linux/kernel/fork.c

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/*
* linux/kernel/fork.c
*
* Copyright (C) 1991, 1992 Linus Torvalds
*/
/*
* 'fork.c' contains the help-routines for the 'fork' system call
* (see also entry.S and others).
* Fork is rather simple, once you get the hang of it, but the memory
* management can be a bitch. See 'mm/memory.c': 'copy_page_range()'
*/
#include <linux/slab.h>
#include <linux/sched/autogroup.h>
#include <linux/sched/mm.h>
#include <linux/sched/coredump.h>
#include <linux/sched/user.h>
#include <linux/sched/numa_balancing.h>
#include <linux/sched/stat.h>
#include <linux/sched/task.h>
#include <linux/sched/task_stack.h>
#include <linux/sched/cputime.h>
#include <linux/rtmutex.h>
#include <linux/init.h>
#include <linux/unistd.h>
#include <linux/module.h>
#include <linux/vmalloc.h>
#include <linux/completion.h>
#include <linux/personality.h>
#include <linux/mempolicy.h>
#include <linux/sem.h>
#include <linux/file.h>
#include <linux/fdtable.h>
#include <linux/iocontext.h>
#include <linux/key.h>
#include <linux/binfmts.h>
#include <linux/mman.h>
mmu-notifiers: core With KVM/GFP/XPMEM there isn't just the primary CPU MMU pointing to pages. There are secondary MMUs (with secondary sptes and secondary tlbs) too. sptes in the kvm case are shadow pagetables, but when I say spte in mmu-notifier context, I mean "secondary pte". In GRU case there's no actual secondary pte and there's only a secondary tlb because the GRU secondary MMU has no knowledge about sptes and every secondary tlb miss event in the MMU always generates a page fault that has to be resolved by the CPU (this is not the case of KVM where the a secondary tlb miss will walk sptes in hardware and it will refill the secondary tlb transparently to software if the corresponding spte is present). The same way zap_page_range has to invalidate the pte before freeing the page, the spte (and secondary tlb) must also be invalidated before any page is freed and reused. Currently we take a page_count pin on every page mapped by sptes, but that means the pages can't be swapped whenever they're mapped by any spte because they're part of the guest working set. Furthermore a spte unmap event can immediately lead to a page to be freed when the pin is released (so requiring the same complex and relatively slow tlb_gather smp safe logic we have in zap_page_range and that can be avoided completely if the spte unmap event doesn't require an unpin of the page previously mapped in the secondary MMU). The mmu notifiers allow kvm/GRU/XPMEM to attach to the tsk->mm and know when the VM is swapping or freeing or doing anything on the primary MMU so that the secondary MMU code can drop sptes before the pages are freed, avoiding all page pinning and allowing 100% reliable swapping of guest physical address space. Furthermore it avoids the code that teardown the mappings of the secondary MMU, to implement a logic like tlb_gather in zap_page_range that would require many IPI to flush other cpu tlbs, for each fixed number of spte unmapped. To make an example: if what happens on the primary MMU is a protection downgrade (from writeable to wrprotect) the secondary MMU mappings will be invalidated, and the next secondary-mmu-page-fault will call get_user_pages and trigger a do_wp_page through get_user_pages if it called get_user_pages with write=1, and it'll re-establishing an updated spte or secondary-tlb-mapping on the copied page. Or it will setup a readonly spte or readonly tlb mapping if it's a guest-read, if it calls get_user_pages with write=0. This is just an example. This allows to map any page pointed by any pte (and in turn visible in the primary CPU MMU), into a secondary MMU (be it a pure tlb like GRU, or an full MMU with both sptes and secondary-tlb like the shadow-pagetable layer with kvm), or a remote DMA in software like XPMEM (hence needing of schedule in XPMEM code to send the invalidate to the remote node, while no need to schedule in kvm/gru as it's an immediate event like invalidating primary-mmu pte). At least for KVM without this patch it's impossible to swap guests reliably. And having this feature and removing the page pin allows several other optimizations that simplify life considerably. Dependencies: 1) mm_take_all_locks() to register the mmu notifier when the whole VM isn't doing anything with "mm". This allows mmu notifier users to keep track if the VM is in the middle of the invalidate_range_begin/end critical section with an atomic counter incraese in range_begin and decreased in range_end. No secondary MMU page fault is allowed to map any spte or secondary tlb reference, while the VM is in the middle of range_begin/end as any page returned by get_user_pages in that critical section could later immediately be freed without any further ->invalidate_page notification (invalidate_range_begin/end works on ranges and ->invalidate_page isn't called immediately before freeing the page). To stop all page freeing and pagetable overwrites the mmap_sem must be taken in write mode and all other anon_vma/i_mmap locks must be taken too. 2) It'd be a waste to add branches in the VM if nobody could possibly run KVM/GRU/XPMEM on the kernel, so mmu notifiers will only enabled if CONFIG_KVM=m/y. In the current kernel kvm won't yet take advantage of mmu notifiers, but this already allows to compile a KVM external module against a kernel with mmu notifiers enabled and from the next pull from kvm.git we'll start using them. And GRU/XPMEM will also be able to continue the development by enabling KVM=m in their config, until they submit all GRU/XPMEM GPLv2 code to the mainline kernel. Then they can also enable MMU_NOTIFIERS in the same way KVM does it (even if KVM=n). This guarantees nobody selects MMU_NOTIFIER=y if KVM and GRU and XPMEM are all =n. The mmu_notifier_register call can fail because mm_take_all_locks may be interrupted by a signal and return -EINTR. Because mmu_notifier_reigster is used when a driver startup, a failure can be gracefully handled. Here an example of the change applied to kvm to register the mmu notifiers. Usually when a driver startups other allocations are required anyway and -ENOMEM failure paths exists already. struct kvm *kvm_arch_create_vm(void) { struct kvm *kvm = kzalloc(sizeof(struct kvm), GFP_KERNEL); + int err; if (!kvm) return ERR_PTR(-ENOMEM); INIT_LIST_HEAD(&kvm->arch.active_mmu_pages); + kvm->arch.mmu_notifier.ops = &kvm_mmu_notifier_ops; + err = mmu_notifier_register(&kvm->arch.mmu_notifier, current->mm); + if (err) { + kfree(kvm); + return ERR_PTR(err); + } + return kvm; } mmu_notifier_unregister returns void and it's reliable. The patch also adds a few needed but missing includes that would prevent kernel to compile after these changes on non-x86 archs (x86 didn't need them by luck). [akpm@linux-foundation.org: coding-style fixes] [akpm@linux-foundation.org: fix mm/filemap_xip.c build] [akpm@linux-foundation.org: fix mm/mmu_notifier.c build] Signed-off-by: Andrea Arcangeli <andrea@qumranet.com> Signed-off-by: Nick Piggin <npiggin@suse.de> Signed-off-by: Christoph Lameter <cl@linux-foundation.org> Cc: Jack Steiner <steiner@sgi.com> Cc: Robin Holt <holt@sgi.com> Cc: Nick Piggin <npiggin@suse.de> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: Kanoj Sarcar <kanojsarcar@yahoo.com> Cc: Roland Dreier <rdreier@cisco.com> Cc: Steve Wise <swise@opengridcomputing.com> Cc: Avi Kivity <avi@qumranet.com> Cc: Hugh Dickins <hugh@veritas.com> Cc: Rusty Russell <rusty@rustcorp.com.au> Cc: Anthony Liguori <aliguori@us.ibm.com> Cc: Chris Wright <chrisw@redhat.com> Cc: Marcelo Tosatti <marcelo@kvack.org> Cc: Eric Dumazet <dada1@cosmosbay.com> Cc: "Paul E. McKenney" <paulmck@us.ibm.com> Cc: Izik Eidus <izike@qumranet.com> Cc: Anthony Liguori <aliguori@us.ibm.com> Cc: Rik van Riel <riel@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-07-28 22:46:29 +00:00
#include <linux/mmu_notifier.h>
#include <linux/fs.h>
mm: per-thread vma caching This patch is a continuation of efforts trying to optimize find_vma(), avoiding potentially expensive rbtree walks to locate a vma upon faults. The original approach (https://lkml.org/lkml/2013/11/1/410), where the largest vma was also cached, ended up being too specific and random, thus further comparison with other approaches were needed. There are two things to consider when dealing with this, the cache hit rate and the latency of find_vma(). Improving the hit-rate does not necessarily translate in finding the vma any faster, as the overhead of any fancy caching schemes can be too high to consider. We currently cache the last used vma for the whole address space, which provides a nice optimization, reducing the total cycles in find_vma() by up to 250%, for workloads with good locality. On the other hand, this simple scheme is pretty much useless for workloads with poor locality. Analyzing ebizzy runs shows that, no matter how many threads are running, the mmap_cache hit rate is less than 2%, and in many situations below 1%. The proposed approach is to replace this scheme with a small per-thread cache, maximizing hit rates at a very low maintenance cost. Invalidations are performed by simply bumping up a 32-bit sequence number. The only expensive operation is in the rare case of a seq number overflow, where all caches that share the same address space are flushed. Upon a miss, the proposed replacement policy is based on the page number that contains the virtual address in question. Concretely, the following results are seen on an 80 core, 8 socket x86-64 box: 1) System bootup: Most programs are single threaded, so the per-thread scheme does improve ~50% hit rate by just adding a few more slots to the cache. +----------------+----------+------------------+ | caching scheme | hit-rate | cycles (billion) | +----------------+----------+------------------+ | baseline | 50.61% | 19.90 | | patched | 73.45% | 13.58 | +----------------+----------+------------------+ 2) Kernel build: This one is already pretty good with the current approach as we're dealing with good locality. +----------------+----------+------------------+ | caching scheme | hit-rate | cycles (billion) | +----------------+----------+------------------+ | baseline | 75.28% | 11.03 | | patched | 88.09% | 9.31 | +----------------+----------+------------------+ 3) Oracle 11g Data Mining (4k pages): Similar to the kernel build workload. +----------------+----------+------------------+ | caching scheme | hit-rate | cycles (billion) | +----------------+----------+------------------+ | baseline | 70.66% | 17.14 | | patched | 91.15% | 12.57 | +----------------+----------+------------------+ 4) Ebizzy: There's a fair amount of variation from run to run, but this approach always shows nearly perfect hit rates, while baseline is just about non-existent. The amounts of cycles can fluctuate between anywhere from ~60 to ~116 for the baseline scheme, but this approach reduces it considerably. For instance, with 80 threads: +----------------+----------+------------------+ | caching scheme | hit-rate | cycles (billion) | +----------------+----------+------------------+ | baseline | 1.06% | 91.54 | | patched | 99.97% | 14.18 | +----------------+----------+------------------+ [akpm@linux-foundation.org: fix nommu build, per Davidlohr] [akpm@linux-foundation.org: document vmacache_valid() logic] [akpm@linux-foundation.org: attempt to untangle header files] [akpm@linux-foundation.org: add vmacache_find() BUG_ON] [hughd@google.com: add vmacache_valid_mm() (from Oleg)] [akpm@linux-foundation.org: coding-style fixes] [akpm@linux-foundation.org: adjust and enhance comments] Signed-off-by: Davidlohr Bueso <davidlohr@hp.com> Reviewed-by: Rik van Riel <riel@redhat.com> Acked-by: Linus Torvalds <torvalds@linux-foundation.org> Reviewed-by: Michel Lespinasse <walken@google.com> Cc: Oleg Nesterov <oleg@redhat.com> Tested-by: Hugh Dickins <hughd@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-04-07 22:37:25 +00:00
#include <linux/mm.h>
#include <linux/vmacache.h>
#include <linux/nsproxy.h>
#include <linux/capability.h>
#include <linux/cpu.h>
#include <linux/cgroup.h>
#include <linux/security.h>
hugetlb: reserve huge pages for reliable MAP_PRIVATE hugetlbfs mappings until fork() This patch reserves huge pages at mmap() time for MAP_PRIVATE mappings in a similar manner to the reservations taken for MAP_SHARED mappings. The reserve count is accounted both globally and on a per-VMA basis for private mappings. This guarantees that a process that successfully calls mmap() will successfully fault all pages in the future unless fork() is called. The characteristics of private mappings of hugetlbfs files behaviour after this patch are; 1. The process calling mmap() is guaranteed to succeed all future faults until it forks(). 2. On fork(), the parent may die due to SIGKILL on writes to the private mapping if enough pages are not available for the COW. For reasonably reliable behaviour in the face of a small huge page pool, children of hugepage-aware processes should not reference the mappings; such as might occur when fork()ing to exec(). 3. On fork(), the child VMAs inherit no reserves. Reads on pages already faulted by the parent will succeed. Successful writes will depend on enough huge pages being free in the pool. 4. Quotas of the hugetlbfs mount are checked at reserve time for the mapper and at fault time otherwise. Before this patch, all reads or writes in the child potentially needs page allocations that can later lead to the death of the parent. This applies to reads and writes of uninstantiated pages as well as COW. After the patch it is only a write to an instantiated page that causes problems. Signed-off-by: Mel Gorman <mel@csn.ul.ie> Acked-by: Adam Litke <agl@us.ibm.com> Cc: Andy Whitcroft <apw@shadowen.org> Cc: William Lee Irwin III <wli@holomorphy.com> Cc: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-07-24 04:27:23 +00:00
#include <linux/hugetlb.h>
seccomp: add system call filtering using BPF [This patch depends on luto@mit.edu's no_new_privs patch: https://lkml.org/lkml/2012/1/30/264 The whole series including Andrew's patches can be found here: https://github.com/redpig/linux/tree/seccomp Complete diff here: https://github.com/redpig/linux/compare/1dc65fed...seccomp ] This patch adds support for seccomp mode 2. Mode 2 introduces the ability for unprivileged processes to install system call filtering policy expressed in terms of a Berkeley Packet Filter (BPF) program. This program will be evaluated in the kernel for each system call the task makes and computes a result based on data in the format of struct seccomp_data. A filter program may be installed by calling: struct sock_fprog fprog = { ... }; ... prctl(PR_SET_SECCOMP, SECCOMP_MODE_FILTER, &fprog); The return value of the filter program determines if the system call is allowed to proceed or denied. If the first filter program installed allows prctl(2) calls, then the above call may be made repeatedly by a task to further reduce its access to the kernel. All attached programs must be evaluated before a system call will be allowed to proceed. Filter programs will be inherited across fork/clone and execve. However, if the task attaching the filter is unprivileged (!CAP_SYS_ADMIN) the no_new_privs bit will be set on the task. This ensures that unprivileged tasks cannot attach filters that affect privileged tasks (e.g., setuid binary). There are a number of benefits to this approach. A few of which are as follows: - BPF has been exposed to userland for a long time - BPF optimization (and JIT'ing) are well understood - Userland already knows its ABI: system call numbers and desired arguments - No time-of-check-time-of-use vulnerable data accesses are possible. - system call arguments are loaded on access only to minimize copying required for system call policy decisions. Mode 2 support is restricted to architectures that enable HAVE_ARCH_SECCOMP_FILTER. In this patch, the primary dependency is on syscall_get_arguments(). The full desired scope of this feature will add a few minor additional requirements expressed later in this series. Based on discussion, SECCOMP_RET_ERRNO and SECCOMP_RET_TRACE seem to be the desired additional functionality. No architectures are enabled in this patch. Signed-off-by: Will Drewry <wad@chromium.org> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Reviewed-by: Indan Zupancic <indan@nul.nu> Acked-by: Eric Paris <eparis@redhat.com> Reviewed-by: Kees Cook <keescook@chromium.org> v18: - rebase to v3.4-rc2 - s/chk/check/ (akpm@linux-foundation.org,jmorris@namei.org) - allocate with GFP_KERNEL|__GFP_NOWARN (indan@nul.nu) - add a comment for get_u32 regarding endianness (akpm@) - fix other typos, style mistakes (akpm@) - added acked-by v17: - properly guard seccomp filter needed headers (leann@ubuntu.com) - tighten return mask to 0x7fff0000 v16: - no change v15: - add a 4 instr penalty when counting a path to account for seccomp_filter size (indan@nul.nu) - drop the max insns to 256KB (indan@nul.nu) - return ENOMEM if the max insns limit has been hit (indan@nul.nu) - move IP checks after args (indan@nul.nu) - drop !user_filter check (indan@nul.nu) - only allow explicit bpf codes (indan@nul.nu) - exit_code -> exit_sig v14: - put/get_seccomp_filter takes struct task_struct (indan@nul.nu,keescook@chromium.org) - adds seccomp_chk_filter and drops general bpf_run/chk_filter user - add seccomp_bpf_load for use by net/core/filter.c - lower max per-process/per-hierarchy: 1MB - moved nnp/capability check prior to allocation (all of the above: indan@nul.nu) v13: - rebase on to 88ebdda6159ffc15699f204c33feb3e431bf9bdc v12: - added a maximum instruction count per path (indan@nul.nu,oleg@redhat.com) - removed copy_seccomp (keescook@chromium.org,indan@nul.nu) - reworded the prctl_set_seccomp comment (indan@nul.nu) v11: - reorder struct seccomp_data to allow future args expansion (hpa@zytor.com) - style clean up, @compat dropped, compat_sock_fprog32 (indan@nul.nu) - do_exit(SIGSYS) (keescook@chromium.org, luto@mit.edu) - pare down Kconfig doc reference. - extra comment clean up v10: - seccomp_data has changed again to be more aesthetically pleasing (hpa@zytor.com) - calling convention is noted in a new u32 field using syscall_get_arch. This allows for cross-calling convention tasks to use seccomp filters. (hpa@zytor.com) - lots of clean up (thanks, Indan!) v9: - n/a v8: - use bpf_chk_filter, bpf_run_filter. update load_fns - Lots of fixes courtesy of indan@nul.nu: -- fix up load behavior, compat fixups, and merge alloc code, -- renamed pc and dropped __packed, use bool compat. -- Added a hidden CONFIG_SECCOMP_FILTER to synthesize non-arch dependencies v7: (massive overhaul thanks to Indan, others) - added CONFIG_HAVE_ARCH_SECCOMP_FILTER - merged into seccomp.c - minimal seccomp_filter.h - no config option (part of seccomp) - no new prctl - doesn't break seccomp on systems without asm/syscall.h (works but arg access always fails) - dropped seccomp_init_task, extra free functions, ... - dropped the no-asm/syscall.h code paths - merges with network sk_run_filter and sk_chk_filter v6: - fix memory leak on attach compat check failure - require no_new_privs || CAP_SYS_ADMIN prior to filter installation. (luto@mit.edu) - s/seccomp_struct_/seccomp_/ for macros/functions (amwang@redhat.com) - cleaned up Kconfig (amwang@redhat.com) - on block, note if the call was compat (so the # means something) v5: - uses syscall_get_arguments (indan@nul.nu,oleg@redhat.com, mcgrathr@chromium.org) - uses union-based arg storage with hi/lo struct to handle endianness. Compromises between the two alternate proposals to minimize extra arg shuffling and account for endianness assuming userspace uses offsetof(). (mcgrathr@chromium.org, indan@nul.nu) - update Kconfig description - add include/seccomp_filter.h and add its installation - (naive) on-demand syscall argument loading - drop seccomp_t (eparis@redhat.com) v4: - adjusted prctl to make room for PR_[SG]ET_NO_NEW_PRIVS - now uses current->no_new_privs (luto@mit.edu,torvalds@linux-foundation.com) - assign names to seccomp modes (rdunlap@xenotime.net) - fix style issues (rdunlap@xenotime.net) - reworded Kconfig entry (rdunlap@xenotime.net) v3: - macros to inline (oleg@redhat.com) - init_task behavior fixed (oleg@redhat.com) - drop creator entry and extra NULL check (oleg@redhat.com) - alloc returns -EINVAL on bad sizing (serge.hallyn@canonical.com) - adds tentative use of "always_unprivileged" as per torvalds@linux-foundation.org and luto@mit.edu v2: - (patch 2 only) Signed-off-by: James Morris <james.l.morris@oracle.com>
2012-04-12 21:47:57 +00:00
#include <linux/seccomp.h>
#include <linux/swap.h>
#include <linux/syscalls.h>
#include <linux/jiffies.h>
#include <linux/futex.h>
#include <linux/compat.h>
#include <linux/kthread.h>
[PATCH] io-accounting: core statistics The present per-task IO accounting isn't very useful. It simply counts the number of bytes passed into read() and write(). So if a process reads 1MB from an already-cached file, it is accused of having performed 1MB of I/O, which is wrong. (David Wright had some comments on the applicability of the present logical IO accounting: For billing purposes it is useless but for workload analysis it is very useful read_bytes/read_calls average read request size write_bytes/write_calls average write request size read_bytes/read_blocks ie logical/physical can indicate hit rate or thrashing write_bytes/write_blocks ie logical/physical guess since pdflush writes can be missed I often look for logical larger than physical to see filesystem cache problems. And the bytes/cpusec can help find applications that are dominating the cache and causing slow interactive response from page cache contention. I want to find the IO intensive applications and make sure they are doing efficient IO. Thus the acctcms(sysV) or csacms command would give the high IO commands). This patchset adds new accounting which tries to be more accurate. We account for three things: reads: attempt to count the number of bytes which this process really did cause to be fetched from the storage layer. Done at the submit_bio() level, so it is accurate for block-backed filesystems. I also attempt to wire up NFS and CIFS. writes: attempt to count the number of bytes which this process caused to be sent to the storage layer. This is done at page-dirtying time. The big inaccuracy here is truncate. If a process writes 1MB to a file and then deletes the file, it will in fact perform no writeout. But it will have been accounted as having caused 1MB of write. So... cancelled_writes: account the number of bytes which this process caused to not happen, by truncating pagecache. We _could_ just subtract this from the process's `write' accounting. But that means that some processes would be reported to have done negative amounts of write IO, which is silly. So we just report the raw number and punt this decision up to userspace. Now, we _could_ account for writes at the physical I/O level. But - This would require that we track memory-dirtying tasks at the per-page level (would require a new pointer in struct page). - It would mean that IO statistics for a process are usually only available long after that process has exitted. Which means that we probably cannot communicate this info via taskstats. This patch: Wire up the kernel-private data structures and the accessor functions to manipulate them. Cc: Jay Lan <jlan@sgi.com> Cc: Shailabh Nagar <nagar@watson.ibm.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Chris Sturtivant <csturtiv@sgi.com> Cc: Tony Ernst <tee@sgi.com> Cc: Guillaume Thouvenin <guillaume.thouvenin@bull.net> Cc: David Wright <daw@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-12-10 10:19:19 +00:00
#include <linux/task_io_accounting_ops.h>
#include <linux/rcupdate.h>
#include <linux/ptrace.h>
#include <linux/mount.h>
#include <linux/audit.h>
#include <linux/memcontrol.h>
#include <linux/ftrace.h>
#include <linux/proc_fs.h>
#include <linux/profile.h>
#include <linux/rmap.h>
ksm: the mm interface to ksm This patch presents the mm interface to a dummy version of ksm.c, for better scrutiny of that interface: the real ksm.c follows later. When CONFIG_KSM is not set, madvise(2) reject MADV_MERGEABLE and MADV_UNMERGEABLE with EINVAL, since that seems more helpful than pretending that they can be serviced. But when CONFIG_KSM=y, accept them even if KSM is not currently running, and even on areas which KSM will not touch (e.g. hugetlb or shared file or special driver mappings). Like other madvices, report ENOMEM despite success if any area in the range is unmapped, and use EAGAIN to report out of memory. Define vma flag VM_MERGEABLE to identify an area on which KSM may try merging pages: leave it to ksm_madvise() to decide whether to set it. Define mm flag MMF_VM_MERGEABLE to identify an mm which might contain VM_MERGEABLE areas, to minimize callouts when forking or exiting. Based upon earlier patches by Chris Wright and Izik Eidus. Signed-off-by: Hugh Dickins <hugh.dickins@tiscali.co.uk> Signed-off-by: Chris Wright <chrisw@redhat.com> Signed-off-by: Izik Eidus <ieidus@redhat.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Rik van Riel <riel@redhat.com> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Lee Schermerhorn <lee.schermerhorn@hp.com> Cc: Avi Kivity <avi@redhat.com> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-22 00:01:57 +00:00
#include <linux/ksm.h>
#include <linux/acct.h>
2017-02-22 23:42:27 +00:00
#include <linux/userfaultfd_k.h>
#include <linux/tsacct_kern.h>
#include <linux/cn_proc.h>
#include <linux/freezer.h>
#include <linux/delayacct.h>
#include <linux/taskstats_kern.h>
#include <linux/random.h>
Audit: add TTY input auditing Add TTY input auditing, used to audit system administrator's actions. This is required by various security standards such as DCID 6/3 and PCI to provide non-repudiation of administrator's actions and to allow a review of past actions if the administrator seems to overstep their duties or if the system becomes misconfigured for unknown reasons. These requirements do not make it necessary to audit TTY output as well. Compared to an user-space keylogger, this approach records TTY input using the audit subsystem, correlated with other audit events, and it is completely transparent to the user-space application (e.g. the console ioctls still work). TTY input auditing works on a higher level than auditing all system calls within the session, which would produce an overwhelming amount of mostly useless audit events. Add an "audit_tty" attribute, inherited across fork (). Data read from TTYs by process with the attribute is sent to the audit subsystem by the kernel. The audit netlink interface is extended to allow modifying the audit_tty attribute, and to allow sending explanatory audit events from user-space (for example, a shell might send an event containing the final command, after the interactive command-line editing and history expansion is performed, which might be difficult to decipher from the TTY input alone). Because the "audit_tty" attribute is inherited across fork (), it would be set e.g. for sshd restarted within an audited session. To prevent this, the audit_tty attribute is cleared when a process with no open TTY file descriptors (e.g. after daemon startup) opens a TTY. See https://www.redhat.com/archives/linux-audit/2007-June/msg00000.html for a more detailed rationale document for an older version of this patch. [akpm@linux-foundation.org: build fix] Signed-off-by: Miloslav Trmac <mitr@redhat.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Alan Cox <alan@lxorguk.ukuu.org.uk> Cc: Paul Fulghum <paulkf@microgate.com> Cc: Casey Schaufler <casey@schaufler-ca.com> Cc: Steve Grubb <sgrubb@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-07-16 06:40:56 +00:00
#include <linux/tty.h>
#include <linux/blkdev.h>
#include <linux/fs_struct.h>
#include <linux/magic.h>
perf: Do the big rename: Performance Counters -> Performance Events Bye-bye Performance Counters, welcome Performance Events! In the past few months the perfcounters subsystem has grown out its initial role of counting hardware events, and has become (and is becoming) a much broader generic event enumeration, reporting, logging, monitoring, analysis facility. Naming its core object 'perf_counter' and naming the subsystem 'perfcounters' has become more and more of a misnomer. With pending code like hw-breakpoints support the 'counter' name is less and less appropriate. All in one, we've decided to rename the subsystem to 'performance events' and to propagate this rename through all fields, variables and API names. (in an ABI compatible fashion) The word 'event' is also a bit shorter than 'counter' - which makes it slightly more convenient to write/handle as well. Thanks goes to Stephane Eranian who first observed this misnomer and suggested a rename. User-space tooling and ABI compatibility is not affected - this patch should be function-invariant. (Also, defconfigs were not touched to keep the size down.) This patch has been generated via the following script: FILES=$(find * -type f | grep -vE 'oprofile|[^K]config') sed -i \ -e 's/PERF_EVENT_/PERF_RECORD_/g' \ -e 's/PERF_COUNTER/PERF_EVENT/g' \ -e 's/perf_counter/perf_event/g' \ -e 's/nb_counters/nb_events/g' \ -e 's/swcounter/swevent/g' \ -e 's/tpcounter_event/tp_event/g' \ $FILES for N in $(find . -name perf_counter.[ch]); do M=$(echo $N | sed 's/perf_counter/perf_event/g') mv $N $M done FILES=$(find . -name perf_event.*) sed -i \ -e 's/COUNTER_MASK/REG_MASK/g' \ -e 's/COUNTER/EVENT/g' \ -e 's/\<event\>/event_id/g' \ -e 's/counter/event/g' \ -e 's/Counter/Event/g' \ $FILES ... to keep it as correct as possible. This script can also be used by anyone who has pending perfcounters patches - it converts a Linux kernel tree over to the new naming. We tried to time this change to the point in time where the amount of pending patches is the smallest: the end of the merge window. Namespace clashes were fixed up in a preparatory patch - and some stylistic fallout will be fixed up in a subsequent patch. ( NOTE: 'counters' are still the proper terminology when we deal with hardware registers - and these sed scripts are a bit over-eager in renaming them. I've undone some of that, but in case there's something left where 'counter' would be better than 'event' we can undo that on an individual basis instead of touching an otherwise nicely automated patch. ) Suggested-by: Stephane Eranian <eranian@google.com> Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl> Acked-by: Paul Mackerras <paulus@samba.org> Reviewed-by: Arjan van de Ven <arjan@linux.intel.com> Cc: Mike Galbraith <efault@gmx.de> Cc: Arnaldo Carvalho de Melo <acme@redhat.com> Cc: Frederic Weisbecker <fweisbec@gmail.com> Cc: Steven Rostedt <rostedt@goodmis.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: David Howells <dhowells@redhat.com> Cc: Kyle McMartin <kyle@mcmartin.ca> Cc: Martin Schwidefsky <schwidefsky@de.ibm.com> Cc: "David S. Miller" <davem@davemloft.net> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: <linux-arch@vger.kernel.org> LKML-Reference: <new-submission> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-09-21 10:02:48 +00:00
#include <linux/perf_event.h>
#include <linux/posix-timers.h>
#include <linux/user-return-notifier.h>
#include <linux/oom.h>
2011-01-13 23:46:58 +00:00
#include <linux/khugepaged.h>
epoll: introduce POLLFREE to flush ->signalfd_wqh before kfree() This patch is intentionally incomplete to simplify the review. It ignores ep_unregister_pollwait() which plays with the same wqh. See the next change. epoll assumes that the EPOLL_CTL_ADD'ed file controls everything f_op->poll() needs. In particular it assumes that the wait queue can't go away until eventpoll_release(). This is not true in case of signalfd, the task which does EPOLL_CTL_ADD uses its ->sighand which is not connected to the file. This patch adds the special event, POLLFREE, currently only for epoll. It expects that init_poll_funcptr()'ed hook should do the necessary cleanup. Perhaps it should be defined as EPOLLFREE in eventpoll. __cleanup_sighand() is changed to do wake_up_poll(POLLFREE) if ->signalfd_wqh is not empty, we add the new signalfd_cleanup() helper. ep_poll_callback(POLLFREE) simply does list_del_init(task_list). This make this poll entry inconsistent, but we don't care. If you share epoll fd which contains our sigfd with another process you should blame yourself. signalfd is "really special". I simply do not know how we can define the "right" semantics if it used with epoll. The main problem is, epoll calls signalfd_poll() once to establish the connection with the wait queue, after that signalfd_poll(NULL) returns the different/inconsistent results depending on who does EPOLL_CTL_MOD/signalfd_read/etc. IOW: apart from sigmask, signalfd has nothing to do with the file, it works with the current thread. In short: this patch is the hack which tries to fix the symptoms. It also assumes that nobody can take tasklist_lock under epoll locks, this seems to be true. Note: - we do not have wake_up_all_poll() but wake_up_poll() is fine, poll/epoll doesn't use WQ_FLAG_EXCLUSIVE. - signalfd_cleanup() uses POLLHUP along with POLLFREE, we need a couple of simple changes in eventpoll.c to make sure it can't be "lost". Reported-by: Maxime Bizon <mbizon@freebox.fr> Cc: <stable@kernel.org> Signed-off-by: Oleg Nesterov <oleg@redhat.com> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-02-24 19:07:11 +00:00
#include <linux/signalfd.h>
uprobes/core: Handle breakpoint and singlestep exceptions Uprobes uses exception notifiers to get to know if a thread hit a breakpoint or a singlestep exception. When a thread hits a uprobe or is singlestepping post a uprobe hit, the uprobe exception notifier sets its TIF_UPROBE bit, which will then be checked on its return to userspace path (do_notify_resume() ->uprobe_notify_resume()), where the consumers handlers are run (in task context) based on the defined filters. Uprobe hits are thread specific and hence we need to maintain information about if a task hit a uprobe, what uprobe was hit, the slot where the original instruction was copied for xol so that it can be singlestepped with appropriate fixups. In some cases, special care is needed for instructions that are executed out of line (xol). These are architecture specific artefacts, such as handling RIP relative instructions on x86_64. Since the instruction at which the uprobe was inserted is executed out of line, architecture specific fixups are added so that the thread continues normal execution in the presence of a uprobe. Postpone the signals until we execute the probed insn. post_xol() path does a recalc_sigpending() before return to user-mode, this ensures the signal can't be lost. Uprobes relies on DIE_DEBUG notification to notify if a singlestep is complete. Adds x86 specific uprobe exception notifiers and appropriate hooks needed to determine a uprobe hit and subsequent post processing. Add requisite x86 fixups for xol for uprobes. Specific cases needing fixups include relative jumps (x86_64), calls, etc. Where possible, we check and skip singlestepping the breakpointed instructions. For now we skip single byte as well as few multibyte nop instructions. However this can be extended to other instructions too. Credits to Oleg Nesterov for suggestions/patches related to signal, breakpoint, singlestep handling code. Signed-off-by: Srikar Dronamraju <srikar@linux.vnet.ibm.com> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Ananth N Mavinakayanahalli <ananth@in.ibm.com> Cc: Jim Keniston <jkenisto@linux.vnet.ibm.com> Cc: Linux-mm <linux-mm@kvack.org> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Andi Kleen <andi@firstfloor.org> Cc: Christoph Hellwig <hch@infradead.org> Cc: Steven Rostedt <rostedt@goodmis.org> Cc: Arnaldo Carvalho de Melo <acme@infradead.org> Cc: Masami Hiramatsu <masami.hiramatsu.pt@hitachi.com> Cc: Peter Zijlstra <peterz@infradead.org> Link: http://lkml.kernel.org/r/20120313180011.29771.89027.sendpatchset@srdronam.in.ibm.com [ Performed various cleanliness edits ] Signed-off-by: Ingo Molnar <mingo@elte.hu>
2012-03-13 18:00:11 +00:00
#include <linux/uprobes.h>
#include <linux/aio.h>
#include <linux/compiler.h>
#include <linux/sysctl.h>
kernel: add kcov code coverage kcov provides code coverage collection for coverage-guided fuzzing (randomized testing). Coverage-guided fuzzing is a testing technique that uses coverage feedback to determine new interesting inputs to a system. A notable user-space example is AFL (http://lcamtuf.coredump.cx/afl/). However, this technique is not widely used for kernel testing due to missing compiler and kernel support. kcov does not aim to collect as much coverage as possible. It aims to collect more or less stable coverage that is function of syscall inputs. To achieve this goal it does not collect coverage in soft/hard interrupts and instrumentation of some inherently non-deterministic or non-interesting parts of kernel is disbled (e.g. scheduler, locking). Currently there is a single coverage collection mode (tracing), but the API anticipates additional collection modes. Initially I also implemented a second mode which exposes coverage in a fixed-size hash table of counters (what Quentin used in his original patch). I've dropped the second mode for simplicity. This patch adds the necessary support on kernel side. The complimentary compiler support was added in gcc revision 231296. We've used this support to build syzkaller system call fuzzer, which has found 90 kernel bugs in just 2 months: https://github.com/google/syzkaller/wiki/Found-Bugs We've also found 30+ bugs in our internal systems with syzkaller. Another (yet unexplored) direction where kcov coverage would greatly help is more traditional "blob mutation". For example, mounting a random blob as a filesystem, or receiving a random blob over wire. Why not gcov. Typical fuzzing loop looks as follows: (1) reset coverage, (2) execute a bit of code, (3) collect coverage, repeat. A typical coverage can be just a dozen of basic blocks (e.g. an invalid input). In such context gcov becomes prohibitively expensive as reset/collect coverage steps depend on total number of basic blocks/edges in program (in case of kernel it is about 2M). Cost of kcov depends only on number of executed basic blocks/edges. On top of that, kernel requires per-thread coverage because there are always background threads and unrelated processes that also produce coverage. With inlined gcov instrumentation per-thread coverage is not possible. kcov exposes kernel PCs and control flow to user-space which is insecure. But debugfs should not be mapped as user accessible. Based on a patch by Quentin Casasnovas. [akpm@linux-foundation.org: make task_struct.kcov_mode have type `enum kcov_mode'] [akpm@linux-foundation.org: unbreak allmodconfig] [akpm@linux-foundation.org: follow x86 Makefile layout standards] Signed-off-by: Dmitry Vyukov <dvyukov@google.com> Reviewed-by: Kees Cook <keescook@chromium.org> Cc: syzkaller <syzkaller@googlegroups.com> Cc: Vegard Nossum <vegard.nossum@oracle.com> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Tavis Ormandy <taviso@google.com> Cc: Will Deacon <will.deacon@arm.com> Cc: Quentin Casasnovas <quentin.casasnovas@oracle.com> Cc: Kostya Serebryany <kcc@google.com> Cc: Eric Dumazet <edumazet@google.com> Cc: Alexander Potapenko <glider@google.com> Cc: Kees Cook <keescook@google.com> Cc: Bjorn Helgaas <bhelgaas@google.com> Cc: Sasha Levin <sasha.levin@oracle.com> Cc: David Drysdale <drysdale@google.com> Cc: Ard Biesheuvel <ard.biesheuvel@linaro.org> Cc: Andrey Ryabinin <ryabinin.a.a@gmail.com> Cc: Kirill A. Shutemov <kirill@shutemov.name> Cc: Jiri Slaby <jslaby@suse.cz> Cc: Ingo Molnar <mingo@elte.hu> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: "H. Peter Anvin" <hpa@zytor.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-03-22 21:27:30 +00:00
#include <linux/kcov.h>
#include <asm/pgtable.h>
#include <asm/pgalloc.h>
#include <linux/uaccess.h>
#include <asm/mmu_context.h>
#include <asm/cacheflush.h>
#include <asm/tlbflush.h>
#include <trace/events/sched.h>
tracepoint: add tracepoints for debugging oom_score_adj oom_score_adj is used for guarding processes from OOM-Killer. One of problem is that it's inherited at fork(). When a daemon set oom_score_adj and make children, it's hard to know where the value is set. This patch adds some tracepoints useful for debugging. This patch adds 3 trace points. - creating new task - renaming a task (exec) - set oom_score_adj To debug, users need to enable some trace pointer. Maybe filtering is useful as # EVENT=/sys/kernel/debug/tracing/events/task/ # echo "oom_score_adj != 0" > $EVENT/task_newtask/filter # echo "oom_score_adj != 0" > $EVENT/task_rename/filter # echo 1 > $EVENT/enable # EVENT=/sys/kernel/debug/tracing/events/oom/ # echo 1 > $EVENT/enable output will be like this. # grep oom /sys/kernel/debug/tracing/trace bash-7699 [007] d..3 5140.744510: oom_score_adj_update: pid=7699 comm=bash oom_score_adj=-1000 bash-7699 [007] ...1 5151.818022: task_newtask: pid=7729 comm=bash clone_flags=1200011 oom_score_adj=-1000 ls-7729 [003] ...2 5151.818504: task_rename: pid=7729 oldcomm=bash newcomm=ls oom_score_adj=-1000 bash-7699 [002] ...1 5175.701468: task_newtask: pid=7730 comm=bash clone_flags=1200011 oom_score_adj=-1000 grep-7730 [007] ...2 5175.701993: task_rename: pid=7730 oldcomm=bash newcomm=grep oom_score_adj=-1000 Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Acked-by: David Rientjes <rientjes@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-01-10 23:08:09 +00:00
#define CREATE_TRACE_POINTS
#include <trace/events/task.h>
/*
* Minimum number of threads to boot the kernel
*/
#define MIN_THREADS 20
/*
* Maximum number of threads
*/
#define MAX_THREADS FUTEX_TID_MASK
/*
* Protected counters by write_lock_irq(&tasklist_lock)
*/
unsigned long total_forks; /* Handle normal Linux uptimes. */
int nr_threads; /* The idle threads do not count.. */
int max_threads; /* tunable limit on nr_threads */
DEFINE_PER_CPU(unsigned long, process_counts) = 0;
__cacheline_aligned DEFINE_RWLOCK(tasklist_lock); /* outer */
#ifdef CONFIG_PROVE_RCU
int lockdep_tasklist_lock_is_held(void)
{
return lockdep_is_held(&tasklist_lock);
}
EXPORT_SYMBOL_GPL(lockdep_tasklist_lock_is_held);
#endif /* #ifdef CONFIG_PROVE_RCU */
int nr_processes(void)
{
int cpu;
int total = 0;
Correct nr_processes() when CPUs have been unplugged nr_processes() returns the sum of the per cpu counter process_counts for all online CPUs. This counter is incremented for the current CPU on fork() and decremented for the current CPU on exit(). Since a process does not necessarily fork and exit on the same CPU the process_count for an individual CPU can be either positive or negative and effectively has no meaning in isolation. Therefore calculating the sum of process_counts over only the online CPUs omits the processes which were started or stopped on any CPU which has since been unplugged. Only the sum of process_counts across all possible CPUs has meaning. The only caller of nr_processes() is proc_root_getattr() which calculates the number of links to /proc as stat->nlink = proc_root.nlink + nr_processes(); You don't have to be all that unlucky for the nr_processes() to return a negative value leading to a negative number of links (or rather, an apparently enormous number of links). If this happens then you can get failures where things like "ls /proc" start to fail because they got an -EOVERFLOW from some stat() call. Example with some debugging inserted to show what goes on: # ps haux|wc -l nr_processes: CPU0: 90 nr_processes: CPU1: 1030 nr_processes: CPU2: -900 nr_processes: CPU3: -136 nr_processes: TOTAL: 84 proc_root_getattr. nlink 12 + nr_processes() 84 = 96 84 # echo 0 >/sys/devices/system/cpu/cpu1/online # ps haux|wc -l nr_processes: CPU0: 85 nr_processes: CPU2: -901 nr_processes: CPU3: -137 nr_processes: TOTAL: -953 proc_root_getattr. nlink 12 + nr_processes() -953 = -941 75 # stat /proc/ nr_processes: CPU0: 84 nr_processes: CPU2: -901 nr_processes: CPU3: -137 nr_processes: TOTAL: -954 proc_root_getattr. nlink 12 + nr_processes() -954 = -942 File: `/proc/' Size: 0 Blocks: 0 IO Block: 1024 directory Device: 3h/3d Inode: 1 Links: 4294966354 Access: (0555/dr-xr-xr-x) Uid: ( 0/ root) Gid: ( 0/ root) Access: 2009-11-03 09:06:55.000000000 +0000 Modify: 2009-11-03 09:06:55.000000000 +0000 Change: 2009-11-03 09:06:55.000000000 +0000 I'm not 100% convinced that the per_cpu regions remain valid for offline CPUs, although my testing suggests that they do. If not then I think the correct solution would be to aggregate the process_count for a given CPU into a global base value in cpu_down(). This bug appears to pre-date the transition to git and it looks like it may even have been present in linux-2.6.0-test7-bk3 since it looks like the code Rusty patched in http://lwn.net/Articles/64773/ was already wrong. Signed-off-by: Ian Campbell <ian.campbell@citrix.com> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Rusty Russell <rusty@rustcorp.com.au> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-11-03 10:11:14 +00:00
for_each_possible_cpu(cpu)
total += per_cpu(process_counts, cpu);
return total;
}
fork: fix error handling in dup_task() The function dup_task() may fail at the following function calls in the following order. 0) alloc_task_struct_node() 1) alloc_thread_info_node() 2) arch_dup_task_struct() Error by 0) is not a matter, it can just return. But error by 1) requires releasing task_struct allocated by 0) before it returns. Likewise, error by 2) requires releasing task_struct and thread_info allocated by 0) and 1). The existing error handling calls free_task_struct() and free_thread_info() which do not only release task_struct and thread_info, but also call architecture specific arch_release_task_struct() and arch_release_thread_info(). The problem is that task_struct and thread_info are not fully initialized yet at this point, but arch_release_task_struct() and arch_release_thread_info() are called with them. For example, x86 defines its own arch_release_task_struct() that releases a task_xstate. If alloc_thread_info_node() fails in dup_task(), arch_release_task_struct() is called with task_struct which is just allocated and filled with garbage in this error handling. This actually happened with tools/testing/fault-injection/failcmd.sh # env FAILCMD_TYPE=fail_page_alloc \ ./tools/testing/fault-injection/failcmd.sh --times=100 \ --min-order=0 --ignore-gfp-wait=0 \ -- make -C tools/testing/selftests/ run_tests In order to fix this issue, make free_{task_struct,thread_info}() not to call arch_release_{task_struct,thread_info}() and call arch_release_{task_struct,thread_info}() implicitly where needed. Default arch_release_task_struct() and arch_release_thread_info() are defined as empty by default. So this change only affects the architectures which implement their own arch_release_task_struct() or arch_release_thread_info() as listed below. arch_release_task_struct(): x86, sh arch_release_thread_info(): mn10300, tile Signed-off-by: Akinobu Mita <akinobu.mita@gmail.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: David Howells <dhowells@redhat.com> Cc: Koichi Yasutake <yasutake.koichi@jp.panasonic.com> Cc: Paul Mundt <lethal@linux-sh.org> Cc: Chris Metcalf <cmetcalf@tilera.com> Cc: Salman Qazi <sqazi@google.com> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: Ingo Molnar <mingo@elte.hu> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-07-30 21:42:33 +00:00
void __weak arch_release_task_struct(struct task_struct *tsk)
{
}
#ifndef CONFIG_ARCH_TASK_STRUCT_ALLOCATOR
static struct kmem_cache *task_struct_cachep;
static inline struct task_struct *alloc_task_struct_node(int node)
{
return kmem_cache_alloc_node(task_struct_cachep, GFP_KERNEL, node);
}
static inline void free_task_struct(struct task_struct *tsk)
{
kmem_cache_free(task_struct_cachep, tsk);
}
#endif
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
void __weak arch_release_thread_stack(unsigned long *stack)
fork: fix error handling in dup_task() The function dup_task() may fail at the following function calls in the following order. 0) alloc_task_struct_node() 1) alloc_thread_info_node() 2) arch_dup_task_struct() Error by 0) is not a matter, it can just return. But error by 1) requires releasing task_struct allocated by 0) before it returns. Likewise, error by 2) requires releasing task_struct and thread_info allocated by 0) and 1). The existing error handling calls free_task_struct() and free_thread_info() which do not only release task_struct and thread_info, but also call architecture specific arch_release_task_struct() and arch_release_thread_info(). The problem is that task_struct and thread_info are not fully initialized yet at this point, but arch_release_task_struct() and arch_release_thread_info() are called with them. For example, x86 defines its own arch_release_task_struct() that releases a task_xstate. If alloc_thread_info_node() fails in dup_task(), arch_release_task_struct() is called with task_struct which is just allocated and filled with garbage in this error handling. This actually happened with tools/testing/fault-injection/failcmd.sh # env FAILCMD_TYPE=fail_page_alloc \ ./tools/testing/fault-injection/failcmd.sh --times=100 \ --min-order=0 --ignore-gfp-wait=0 \ -- make -C tools/testing/selftests/ run_tests In order to fix this issue, make free_{task_struct,thread_info}() not to call arch_release_{task_struct,thread_info}() and call arch_release_{task_struct,thread_info}() implicitly where needed. Default arch_release_task_struct() and arch_release_thread_info() are defined as empty by default. So this change only affects the architectures which implement their own arch_release_task_struct() or arch_release_thread_info() as listed below. arch_release_task_struct(): x86, sh arch_release_thread_info(): mn10300, tile Signed-off-by: Akinobu Mita <akinobu.mita@gmail.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: David Howells <dhowells@redhat.com> Cc: Koichi Yasutake <yasutake.koichi@jp.panasonic.com> Cc: Paul Mundt <lethal@linux-sh.org> Cc: Chris Metcalf <cmetcalf@tilera.com> Cc: Salman Qazi <sqazi@google.com> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: Ingo Molnar <mingo@elte.hu> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-07-30 21:42:33 +00:00
{
}
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
#ifndef CONFIG_ARCH_THREAD_STACK_ALLOCATOR
/*
* Allocate pages if THREAD_SIZE is >= PAGE_SIZE, otherwise use a
* kmemcache based allocator.
*/
# if THREAD_SIZE >= PAGE_SIZE || defined(CONFIG_VMAP_STACK)
#ifdef CONFIG_VMAP_STACK
/*
* vmalloc() is a bit slow, and calling vfree() enough times will force a TLB
* flush. Try to minimize the number of calls by caching stacks.
*/
#define NR_CACHED_STACKS 2
static DEFINE_PER_CPU(struct vm_struct *, cached_stacks[NR_CACHED_STACKS]);
#endif
static unsigned long *alloc_thread_stack_node(struct task_struct *tsk, int node)
{
#ifdef CONFIG_VMAP_STACK
void *stack;
int i;
local_irq_disable();
for (i = 0; i < NR_CACHED_STACKS; i++) {
struct vm_struct *s = this_cpu_read(cached_stacks[i]);
if (!s)
continue;
this_cpu_write(cached_stacks[i], NULL);
tsk->stack_vm_area = s;
local_irq_enable();
return s->addr;
}
local_irq_enable();
stack = __vmalloc_node_range(THREAD_SIZE, THREAD_SIZE,
VMALLOC_START, VMALLOC_END,
THREADINFO_GFP | __GFP_HIGHMEM,
PAGE_KERNEL,
0, node, __builtin_return_address(0));
/*
* We can't call find_vm_area() in interrupt context, and
* free_thread_stack() can be called in interrupt context,
* so cache the vm_struct.
*/
if (stack)
tsk->stack_vm_area = find_vm_area(stack);
return stack;
#else
mm: charge/uncharge kmemcg from generic page allocator paths Currently, to charge a non-slab allocation to kmemcg one has to use alloc_kmem_pages helper with __GFP_ACCOUNT flag. A page allocated with this helper should finally be freed using free_kmem_pages, otherwise it won't be uncharged. This API suits its current users fine, but it turns out to be impossible to use along with page reference counting, i.e. when an allocation is supposed to be freed with put_page, as it is the case with pipe or unix socket buffers. To overcome this limitation, this patch moves charging/uncharging to generic page allocator paths, i.e. to __alloc_pages_nodemask and free_pages_prepare, and zaps alloc/free_kmem_pages helpers. This way, one can use any of the available page allocation functions to get the allocated page charged to kmemcg - it's enough to pass __GFP_ACCOUNT, just like in case of kmalloc and friends. A charged page will be automatically uncharged on free. To make it possible, we need to mark pages charged to kmemcg somehow. To avoid introducing a new page flag, we make use of page->_mapcount for marking such pages. Since pages charged to kmemcg are not supposed to be mapped to userspace, it should work just fine. There are other (ab)users of page->_mapcount - buddy and balloon pages - but we don't conflict with them. In case kmemcg is compiled out or not used at runtime, this patch introduces no overhead to generic page allocator paths. If kmemcg is used, it will be plus one gfp flags check on alloc and plus one page->_mapcount check on free, which shouldn't hurt performance, because the data accessed are hot. Link: http://lkml.kernel.org/r/a9736d856f895bcb465d9f257b54efe32eda6f99.1464079538.git.vdavydov@virtuozzo.com Signed-off-by: Vladimir Davydov <vdavydov@virtuozzo.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@kernel.org> Cc: Eric Dumazet <eric.dumazet@gmail.com> Cc: Minchan Kim <minchan@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-26 22:24:24 +00:00
struct page *page = alloc_pages_node(node, THREADINFO_GFP,
THREAD_SIZE_ORDER);
return page ? page_address(page) : NULL;
#endif
}
static inline void free_thread_stack(struct task_struct *tsk)
{
#ifdef CONFIG_VMAP_STACK
if (task_stack_vm_area(tsk)) {
unsigned long flags;
int i;
local_irq_save(flags);
for (i = 0; i < NR_CACHED_STACKS; i++) {
if (this_cpu_read(cached_stacks[i]))
continue;
this_cpu_write(cached_stacks[i], tsk->stack_vm_area);
local_irq_restore(flags);
return;
}
local_irq_restore(flags);
vfree_atomic(tsk->stack);
return;
}
#endif
__free_pages(virt_to_page(tsk->stack), THREAD_SIZE_ORDER);
}
# else
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
static struct kmem_cache *thread_stack_cache;
static unsigned long *alloc_thread_stack_node(struct task_struct *tsk,
int node)
{
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
return kmem_cache_alloc_node(thread_stack_cache, THREADINFO_GFP, node);
}
static void free_thread_stack(struct task_struct *tsk)
{
kmem_cache_free(thread_stack_cache, tsk->stack);
}
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
void thread_stack_cache_init(void)
{
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
thread_stack_cache = kmem_cache_create("thread_stack", THREAD_SIZE,
THREAD_SIZE, 0, NULL);
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
BUG_ON(thread_stack_cache == NULL);
}
# endif
#endif
/* SLAB cache for signal_struct structures (tsk->signal) */
static struct kmem_cache *signal_cachep;
/* SLAB cache for sighand_struct structures (tsk->sighand) */
struct kmem_cache *sighand_cachep;
/* SLAB cache for files_struct structures (tsk->files) */
struct kmem_cache *files_cachep;
/* SLAB cache for fs_struct structures (tsk->fs) */
struct kmem_cache *fs_cachep;
/* SLAB cache for vm_area_struct structures */
struct kmem_cache *vm_area_cachep;
/* SLAB cache for mm_struct structures (tsk->mm) */
static struct kmem_cache *mm_cachep;
static void account_kernel_stack(struct task_struct *tsk, int account)
{
void *stack = task_stack_page(tsk);
struct vm_struct *vm = task_stack_vm_area(tsk);
BUILD_BUG_ON(IS_ENABLED(CONFIG_VMAP_STACK) && PAGE_SIZE % 1024 != 0);
if (vm) {
int i;
BUG_ON(vm->nr_pages != THREAD_SIZE / PAGE_SIZE);
for (i = 0; i < THREAD_SIZE / PAGE_SIZE; i++) {
mod_zone_page_state(page_zone(vm->pages[i]),
NR_KERNEL_STACK_KB,
PAGE_SIZE / 1024 * account);
}
/* All stack pages belong to the same memcg. */
memcg_kmem_update_page_stat(vm->pages[0], MEMCG_KERNEL_STACK_KB,
account * (THREAD_SIZE / 1024));
} else {
/*
* All stack pages are in the same zone and belong to the
* same memcg.
*/
struct page *first_page = virt_to_page(stack);
mod_zone_page_state(page_zone(first_page), NR_KERNEL_STACK_KB,
THREAD_SIZE / 1024 * account);
memcg_kmem_update_page_stat(first_page, MEMCG_KERNEL_STACK_KB,
account * (THREAD_SIZE / 1024));
}
}
static void release_task_stack(struct task_struct *tsk)
{
if (WARN_ON(tsk->state != TASK_DEAD))
return; /* Better to leak the stack than to free prematurely */
account_kernel_stack(tsk, -1);
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
arch_release_thread_stack(tsk->stack);
free_thread_stack(tsk);
tsk->stack = NULL;
#ifdef CONFIG_VMAP_STACK
tsk->stack_vm_area = NULL;
#endif
}
#ifdef CONFIG_THREAD_INFO_IN_TASK
void put_task_stack(struct task_struct *tsk)
{
if (atomic_dec_and_test(&tsk->stack_refcount))
release_task_stack(tsk);
}
#endif
void free_task(struct task_struct *tsk)
{
#ifndef CONFIG_THREAD_INFO_IN_TASK
/*
* The task is finally done with both the stack and thread_info,
* so free both.
*/
release_task_stack(tsk);
#else
/*
* If the task had a separate stack allocation, it should be gone
* by now.
*/
WARN_ON_ONCE(atomic_read(&tsk->stack_refcount) != 0);
#endif
rt_mutex_debug_task_free(tsk);
ftrace_graph_exit_task(tsk);
seccomp: add system call filtering using BPF [This patch depends on luto@mit.edu's no_new_privs patch: https://lkml.org/lkml/2012/1/30/264 The whole series including Andrew's patches can be found here: https://github.com/redpig/linux/tree/seccomp Complete diff here: https://github.com/redpig/linux/compare/1dc65fed...seccomp ] This patch adds support for seccomp mode 2. Mode 2 introduces the ability for unprivileged processes to install system call filtering policy expressed in terms of a Berkeley Packet Filter (BPF) program. This program will be evaluated in the kernel for each system call the task makes and computes a result based on data in the format of struct seccomp_data. A filter program may be installed by calling: struct sock_fprog fprog = { ... }; ... prctl(PR_SET_SECCOMP, SECCOMP_MODE_FILTER, &fprog); The return value of the filter program determines if the system call is allowed to proceed or denied. If the first filter program installed allows prctl(2) calls, then the above call may be made repeatedly by a task to further reduce its access to the kernel. All attached programs must be evaluated before a system call will be allowed to proceed. Filter programs will be inherited across fork/clone and execve. However, if the task attaching the filter is unprivileged (!CAP_SYS_ADMIN) the no_new_privs bit will be set on the task. This ensures that unprivileged tasks cannot attach filters that affect privileged tasks (e.g., setuid binary). There are a number of benefits to this approach. A few of which are as follows: - BPF has been exposed to userland for a long time - BPF optimization (and JIT'ing) are well understood - Userland already knows its ABI: system call numbers and desired arguments - No time-of-check-time-of-use vulnerable data accesses are possible. - system call arguments are loaded on access only to minimize copying required for system call policy decisions. Mode 2 support is restricted to architectures that enable HAVE_ARCH_SECCOMP_FILTER. In this patch, the primary dependency is on syscall_get_arguments(). The full desired scope of this feature will add a few minor additional requirements expressed later in this series. Based on discussion, SECCOMP_RET_ERRNO and SECCOMP_RET_TRACE seem to be the desired additional functionality. No architectures are enabled in this patch. Signed-off-by: Will Drewry <wad@chromium.org> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Reviewed-by: Indan Zupancic <indan@nul.nu> Acked-by: Eric Paris <eparis@redhat.com> Reviewed-by: Kees Cook <keescook@chromium.org> v18: - rebase to v3.4-rc2 - s/chk/check/ (akpm@linux-foundation.org,jmorris@namei.org) - allocate with GFP_KERNEL|__GFP_NOWARN (indan@nul.nu) - add a comment for get_u32 regarding endianness (akpm@) - fix other typos, style mistakes (akpm@) - added acked-by v17: - properly guard seccomp filter needed headers (leann@ubuntu.com) - tighten return mask to 0x7fff0000 v16: - no change v15: - add a 4 instr penalty when counting a path to account for seccomp_filter size (indan@nul.nu) - drop the max insns to 256KB (indan@nul.nu) - return ENOMEM if the max insns limit has been hit (indan@nul.nu) - move IP checks after args (indan@nul.nu) - drop !user_filter check (indan@nul.nu) - only allow explicit bpf codes (indan@nul.nu) - exit_code -> exit_sig v14: - put/get_seccomp_filter takes struct task_struct (indan@nul.nu,keescook@chromium.org) - adds seccomp_chk_filter and drops general bpf_run/chk_filter user - add seccomp_bpf_load for use by net/core/filter.c - lower max per-process/per-hierarchy: 1MB - moved nnp/capability check prior to allocation (all of the above: indan@nul.nu) v13: - rebase on to 88ebdda6159ffc15699f204c33feb3e431bf9bdc v12: - added a maximum instruction count per path (indan@nul.nu,oleg@redhat.com) - removed copy_seccomp (keescook@chromium.org,indan@nul.nu) - reworded the prctl_set_seccomp comment (indan@nul.nu) v11: - reorder struct seccomp_data to allow future args expansion (hpa@zytor.com) - style clean up, @compat dropped, compat_sock_fprog32 (indan@nul.nu) - do_exit(SIGSYS) (keescook@chromium.org, luto@mit.edu) - pare down Kconfig doc reference. - extra comment clean up v10: - seccomp_data has changed again to be more aesthetically pleasing (hpa@zytor.com) - calling convention is noted in a new u32 field using syscall_get_arch. This allows for cross-calling convention tasks to use seccomp filters. (hpa@zytor.com) - lots of clean up (thanks, Indan!) v9: - n/a v8: - use bpf_chk_filter, bpf_run_filter. update load_fns - Lots of fixes courtesy of indan@nul.nu: -- fix up load behavior, compat fixups, and merge alloc code, -- renamed pc and dropped __packed, use bool compat. -- Added a hidden CONFIG_SECCOMP_FILTER to synthesize non-arch dependencies v7: (massive overhaul thanks to Indan, others) - added CONFIG_HAVE_ARCH_SECCOMP_FILTER - merged into seccomp.c - minimal seccomp_filter.h - no config option (part of seccomp) - no new prctl - doesn't break seccomp on systems without asm/syscall.h (works but arg access always fails) - dropped seccomp_init_task, extra free functions, ... - dropped the no-asm/syscall.h code paths - merges with network sk_run_filter and sk_chk_filter v6: - fix memory leak on attach compat check failure - require no_new_privs || CAP_SYS_ADMIN prior to filter installation. (luto@mit.edu) - s/seccomp_struct_/seccomp_/ for macros/functions (amwang@redhat.com) - cleaned up Kconfig (amwang@redhat.com) - on block, note if the call was compat (so the # means something) v5: - uses syscall_get_arguments (indan@nul.nu,oleg@redhat.com, mcgrathr@chromium.org) - uses union-based arg storage with hi/lo struct to handle endianness. Compromises between the two alternate proposals to minimize extra arg shuffling and account for endianness assuming userspace uses offsetof(). (mcgrathr@chromium.org, indan@nul.nu) - update Kconfig description - add include/seccomp_filter.h and add its installation - (naive) on-demand syscall argument loading - drop seccomp_t (eparis@redhat.com) v4: - adjusted prctl to make room for PR_[SG]ET_NO_NEW_PRIVS - now uses current->no_new_privs (luto@mit.edu,torvalds@linux-foundation.com) - assign names to seccomp modes (rdunlap@xenotime.net) - fix style issues (rdunlap@xenotime.net) - reworded Kconfig entry (rdunlap@xenotime.net) v3: - macros to inline (oleg@redhat.com) - init_task behavior fixed (oleg@redhat.com) - drop creator entry and extra NULL check (oleg@redhat.com) - alloc returns -EINVAL on bad sizing (serge.hallyn@canonical.com) - adds tentative use of "always_unprivileged" as per torvalds@linux-foundation.org and luto@mit.edu v2: - (patch 2 only) Signed-off-by: James Morris <james.l.morris@oracle.com>
2012-04-12 21:47:57 +00:00
put_seccomp_filter(tsk);
fork: fix error handling in dup_task() The function dup_task() may fail at the following function calls in the following order. 0) alloc_task_struct_node() 1) alloc_thread_info_node() 2) arch_dup_task_struct() Error by 0) is not a matter, it can just return. But error by 1) requires releasing task_struct allocated by 0) before it returns. Likewise, error by 2) requires releasing task_struct and thread_info allocated by 0) and 1). The existing error handling calls free_task_struct() and free_thread_info() which do not only release task_struct and thread_info, but also call architecture specific arch_release_task_struct() and arch_release_thread_info(). The problem is that task_struct and thread_info are not fully initialized yet at this point, but arch_release_task_struct() and arch_release_thread_info() are called with them. For example, x86 defines its own arch_release_task_struct() that releases a task_xstate. If alloc_thread_info_node() fails in dup_task(), arch_release_task_struct() is called with task_struct which is just allocated and filled with garbage in this error handling. This actually happened with tools/testing/fault-injection/failcmd.sh # env FAILCMD_TYPE=fail_page_alloc \ ./tools/testing/fault-injection/failcmd.sh --times=100 \ --min-order=0 --ignore-gfp-wait=0 \ -- make -C tools/testing/selftests/ run_tests In order to fix this issue, make free_{task_struct,thread_info}() not to call arch_release_{task_struct,thread_info}() and call arch_release_{task_struct,thread_info}() implicitly where needed. Default arch_release_task_struct() and arch_release_thread_info() are defined as empty by default. So this change only affects the architectures which implement their own arch_release_task_struct() or arch_release_thread_info() as listed below. arch_release_task_struct(): x86, sh arch_release_thread_info(): mn10300, tile Signed-off-by: Akinobu Mita <akinobu.mita@gmail.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: David Howells <dhowells@redhat.com> Cc: Koichi Yasutake <yasutake.koichi@jp.panasonic.com> Cc: Paul Mundt <lethal@linux-sh.org> Cc: Chris Metcalf <cmetcalf@tilera.com> Cc: Salman Qazi <sqazi@google.com> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: Ingo Molnar <mingo@elte.hu> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-07-30 21:42:33 +00:00
arch_release_task_struct(tsk);
if (tsk->flags & PF_KTHREAD)
free_kthread_struct(tsk);
free_task_struct(tsk);
}
EXPORT_SYMBOL(free_task);
static inline void free_signal_struct(struct signal_struct *sig)
{
taskstats_tgid_free(sig);
sched_autogroup_exit(sig);
kernel, oom: fix potential pgd_lock deadlock from __mmdrop Lockdep complains that __mmdrop is not safe from the softirq context: ================================= [ INFO: inconsistent lock state ] 4.6.0-oomfortification2-00011-geeb3eadeab96-dirty #949 Tainted: G W --------------------------------- inconsistent {SOFTIRQ-ON-W} -> {IN-SOFTIRQ-W} usage. swapper/1/0 [HC0[0]:SC1[1]:HE1:SE0] takes: (pgd_lock){+.?...}, at: pgd_free+0x19/0x6b {SOFTIRQ-ON-W} state was registered at: __lock_acquire+0xa06/0x196e lock_acquire+0x139/0x1e1 _raw_spin_lock+0x32/0x41 __change_page_attr_set_clr+0x2a5/0xacd change_page_attr_set_clr+0x16f/0x32c set_memory_nx+0x37/0x3a free_init_pages+0x9e/0xc7 alternative_instructions+0xa2/0xb3 check_bugs+0xe/0x2d start_kernel+0x3ce/0x3ea x86_64_start_reservations+0x2a/0x2c x86_64_start_kernel+0x17a/0x18d irq event stamp: 105916 hardirqs last enabled at (105916): free_hot_cold_page+0x37e/0x390 hardirqs last disabled at (105915): free_hot_cold_page+0x2c1/0x390 softirqs last enabled at (105878): _local_bh_enable+0x42/0x44 softirqs last disabled at (105879): irq_exit+0x6f/0xd1 other info that might help us debug this: Possible unsafe locking scenario: CPU0 ---- lock(pgd_lock); <Interrupt> lock(pgd_lock); *** DEADLOCK *** 1 lock held by swapper/1/0: #0: (rcu_callback){......}, at: rcu_process_callbacks+0x390/0x800 stack backtrace: CPU: 1 PID: 0 Comm: swapper/1 Tainted: G W 4.6.0-oomfortification2-00011-geeb3eadeab96-dirty #949 Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS Debian-1.8.2-1 04/01/2014 Call Trace: <IRQ> print_usage_bug.part.25+0x259/0x268 mark_lock+0x381/0x567 __lock_acquire+0x993/0x196e lock_acquire+0x139/0x1e1 _raw_spin_lock+0x32/0x41 pgd_free+0x19/0x6b __mmdrop+0x25/0xb9 __put_task_struct+0x103/0x11e delayed_put_task_struct+0x157/0x15e rcu_process_callbacks+0x660/0x800 __do_softirq+0x1ec/0x4d5 irq_exit+0x6f/0xd1 smp_apic_timer_interrupt+0x42/0x4d apic_timer_interrupt+0x8e/0xa0 <EOI> arch_cpu_idle+0xf/0x11 default_idle_call+0x32/0x34 cpu_startup_entry+0x20c/0x399 start_secondary+0xfe/0x101 More over commit a79e53d85683 ("x86/mm: Fix pgd_lock deadlock") was explicit about pgd_lock not to be called from the irq context. This means that __mmdrop called from free_signal_struct has to be postponed to a user context. We already have a similar mechanism for mmput_async so we can use it here as well. This is safe because mm_count is pinned by mm_users. This fixes bug introduced by "oom: keep mm of the killed task available" Link: http://lkml.kernel.org/r/1472119394-11342-5-git-send-email-mhocko@kernel.org Signed-off-by: Michal Hocko <mhocko@suse.com> Cc: Tetsuo Handa <penguin-kernel@I-love.SAKURA.ne.jp> Cc: Oleg Nesterov <oleg@redhat.com> Cc: David Rientjes <rientjes@google.com> Cc: Vladimir Davydov <vdavydov@parallels.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-10-07 23:58:54 +00:00
/*
* __mmdrop is not safe to call from softirq context on x86 due to
* pgd_dtor so postpone it to the async context
*/
oom: keep mm of the killed task available oom_reap_task has to call exit_oom_victim in order to make sure that the oom vicim will not block the oom killer for ever. This is, however, opening new problems (e.g oom_killer_disable exclusion - see commit 74070542099c ("oom, suspend: fix oom_reaper vs. oom_killer_disable race")). exit_oom_victim should be only called from the victim's context ideally. One way to achieve this would be to rely on per mm_struct flags. We already have MMF_OOM_REAPED to hide a task from the oom killer since "mm, oom: hide mm which is shared with kthread or global init". The problem is that the exit path: do_exit exit_mm tsk->mm = NULL; mmput __mmput exit_oom_victim doesn't guarantee that exit_oom_victim will get called in a bounded amount of time. At least exit_aio depends on IO which might get blocked due to lack of memory and who knows what else is lurking there. This patch takes a different approach. We remember tsk->mm into the signal_struct and bind it to the signal struct life time for all oom victims. __oom_reap_task_mm as well as oom_scan_process_thread do not have to rely on find_lock_task_mm anymore and they will have a reliable reference to the mm struct. As a result all the oom specific communication inside the OOM killer can be done via tsk->signal->oom_mm. Increasing the signal_struct for something as unlikely as the oom killer is far from ideal but this approach will make the code much more reasonable and long term we even might want to move task->mm into the signal_struct anyway. In the next step we might want to make the oom killer exclusion and access to memory reserves completely independent which would be also nice. Link: http://lkml.kernel.org/r/1472119394-11342-4-git-send-email-mhocko@kernel.org Signed-off-by: Michal Hocko <mhocko@suse.com> Cc: Tetsuo Handa <penguin-kernel@I-love.SAKURA.ne.jp> Cc: Oleg Nesterov <oleg@redhat.com> Cc: David Rientjes <rientjes@google.com> Cc: Vladimir Davydov <vdavydov@parallels.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-10-07 23:58:51 +00:00
if (sig->oom_mm)
kernel, oom: fix potential pgd_lock deadlock from __mmdrop Lockdep complains that __mmdrop is not safe from the softirq context: ================================= [ INFO: inconsistent lock state ] 4.6.0-oomfortification2-00011-geeb3eadeab96-dirty #949 Tainted: G W --------------------------------- inconsistent {SOFTIRQ-ON-W} -> {IN-SOFTIRQ-W} usage. swapper/1/0 [HC0[0]:SC1[1]:HE1:SE0] takes: (pgd_lock){+.?...}, at: pgd_free+0x19/0x6b {SOFTIRQ-ON-W} state was registered at: __lock_acquire+0xa06/0x196e lock_acquire+0x139/0x1e1 _raw_spin_lock+0x32/0x41 __change_page_attr_set_clr+0x2a5/0xacd change_page_attr_set_clr+0x16f/0x32c set_memory_nx+0x37/0x3a free_init_pages+0x9e/0xc7 alternative_instructions+0xa2/0xb3 check_bugs+0xe/0x2d start_kernel+0x3ce/0x3ea x86_64_start_reservations+0x2a/0x2c x86_64_start_kernel+0x17a/0x18d irq event stamp: 105916 hardirqs last enabled at (105916): free_hot_cold_page+0x37e/0x390 hardirqs last disabled at (105915): free_hot_cold_page+0x2c1/0x390 softirqs last enabled at (105878): _local_bh_enable+0x42/0x44 softirqs last disabled at (105879): irq_exit+0x6f/0xd1 other info that might help us debug this: Possible unsafe locking scenario: CPU0 ---- lock(pgd_lock); <Interrupt> lock(pgd_lock); *** DEADLOCK *** 1 lock held by swapper/1/0: #0: (rcu_callback){......}, at: rcu_process_callbacks+0x390/0x800 stack backtrace: CPU: 1 PID: 0 Comm: swapper/1 Tainted: G W 4.6.0-oomfortification2-00011-geeb3eadeab96-dirty #949 Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS Debian-1.8.2-1 04/01/2014 Call Trace: <IRQ> print_usage_bug.part.25+0x259/0x268 mark_lock+0x381/0x567 __lock_acquire+0x993/0x196e lock_acquire+0x139/0x1e1 _raw_spin_lock+0x32/0x41 pgd_free+0x19/0x6b __mmdrop+0x25/0xb9 __put_task_struct+0x103/0x11e delayed_put_task_struct+0x157/0x15e rcu_process_callbacks+0x660/0x800 __do_softirq+0x1ec/0x4d5 irq_exit+0x6f/0xd1 smp_apic_timer_interrupt+0x42/0x4d apic_timer_interrupt+0x8e/0xa0 <EOI> arch_cpu_idle+0xf/0x11 default_idle_call+0x32/0x34 cpu_startup_entry+0x20c/0x399 start_secondary+0xfe/0x101 More over commit a79e53d85683 ("x86/mm: Fix pgd_lock deadlock") was explicit about pgd_lock not to be called from the irq context. This means that __mmdrop called from free_signal_struct has to be postponed to a user context. We already have a similar mechanism for mmput_async so we can use it here as well. This is safe because mm_count is pinned by mm_users. This fixes bug introduced by "oom: keep mm of the killed task available" Link: http://lkml.kernel.org/r/1472119394-11342-5-git-send-email-mhocko@kernel.org Signed-off-by: Michal Hocko <mhocko@suse.com> Cc: Tetsuo Handa <penguin-kernel@I-love.SAKURA.ne.jp> Cc: Oleg Nesterov <oleg@redhat.com> Cc: David Rientjes <rientjes@google.com> Cc: Vladimir Davydov <vdavydov@parallels.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-10-07 23:58:54 +00:00
mmdrop_async(sig->oom_mm);
kmem_cache_free(signal_cachep, sig);
}
static inline void put_signal_struct(struct signal_struct *sig)
{
if (atomic_dec_and_test(&sig->sigcnt))
free_signal_struct(sig);
}
void __put_task_struct(struct task_struct *tsk)
{
WARN_ON(!tsk->exit_state);
WARN_ON(atomic_read(&tsk->usage));
WARN_ON(tsk == current);
cgroup: keep zombies associated with their original cgroups cgroup_exit() is called when a task exits and disassociates the exiting task from its cgroups and half-attach it to the root cgroup. This is unnecessary and undesirable. No controller actually needs an exiting task to be disassociated with non-root cgroups. Both cpu and perf_event controllers update the association to the root cgroup from their exit callbacks just to keep consistent with the cgroup core behavior. Also, this disassociation makes it difficult to track resources held by zombies or determine where the zombies came from. Currently, pids controller is completely broken as it uncharges on exit and zombies always escape the resource restriction. With cgroup association being reset on exit, fixing it is pretty painful. There's no reason to reset cgroup membership on exit. The zombie can be removed from its css_set so that it doesn't show up on "cgroup.procs" and thus can't be migrated or interfere with cgroup removal. It can still pin and point to the css_set so that its cgroup membership is maintained. This patch makes cgroup core keep zombies associated with their cgroups at the time of exit. * Previous patches decoupled populated_cnt tracking from css_set lifetime, so a dying task can be simply unlinked from its css_set while pinning and pointing to the css_set. This keeps css_set association from task side alive while hiding it from "cgroup.procs" and populated_cnt tracking. The css_set reference is dropped when the task_struct is freed. * ->exit() callback no longer needs the css arguments as the associated css never changes once PF_EXITING is set. Removed. * cpu and perf_events controllers no longer need ->exit() callbacks. There's no reason to explicitly switch away on exit. The final schedule out is enough. The callbacks are removed. * On traditional hierarchies, nothing changes. "/proc/PID/cgroup" still reports "/" for all zombies. On the default hierarchy, "/proc/PID/cgroup" keeps reporting the cgroup that the task belonged to at the time of exit. If the cgroup gets removed before the task is reaped, " (deleted)" is appended. v2: Build brekage due to missing dummy cgroup_free() when !CONFIG_CGROUP fixed. Signed-off-by: Tejun Heo <tj@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Arnaldo Carvalho de Melo <acme@kernel.org>
2015-10-15 20:41:53 +00:00
cgroup_free(tsk);
task_numa_free(tsk);
security_task_free(tsk);
exit_creds(tsk);
delayacct_tsk_free(tsk);
put_signal_struct(tsk->signal);
if (!profile_handoff_task(tsk))
free_task(tsk);
}
EXPORT_SYMBOL_GPL(__put_task_struct);
void __init __weak arch_task_cache_init(void) { }
/*
* set_max_threads
*/
static void set_max_threads(unsigned int max_threads_suggested)
{
u64 threads;
/*
* The number of threads shall be limited such that the thread
* structures may only consume a small part of the available memory.
*/
if (fls64(totalram_pages) + fls64(PAGE_SIZE) > 64)
threads = MAX_THREADS;
else
threads = div64_u64((u64) totalram_pages * (u64) PAGE_SIZE,
(u64) THREAD_SIZE * 8UL);
if (threads > max_threads_suggested)
threads = max_threads_suggested;
max_threads = clamp_t(u64, threads, MIN_THREADS, MAX_THREADS);
}
#ifdef CONFIG_ARCH_WANTS_DYNAMIC_TASK_STRUCT
/* Initialized by the architecture: */
int arch_task_struct_size __read_mostly;
#endif
void __init fork_init(void)
{
int i;
#ifndef CONFIG_ARCH_TASK_STRUCT_ALLOCATOR
#ifndef ARCH_MIN_TASKALIGN
#define ARCH_MIN_TASKALIGN 0
#endif
int align = max_t(int, L1_CACHE_BYTES, ARCH_MIN_TASKALIGN);
/* create a slab on which task_structs can be allocated */
2016-01-14 23:18:21 +00:00
task_struct_cachep = kmem_cache_create("task_struct",
arch_task_struct_size, align,
2016-01-14 23:18:21 +00:00
SLAB_PANIC|SLAB_NOTRACK|SLAB_ACCOUNT, NULL);
#endif
/* do the arch specific task caches init */
arch_task_cache_init();
set_max_threads(MAX_THREADS);
init_task.signal->rlim[RLIMIT_NPROC].rlim_cur = max_threads/2;
init_task.signal->rlim[RLIMIT_NPROC].rlim_max = max_threads/2;
init_task.signal->rlim[RLIMIT_SIGPENDING] =
init_task.signal->rlim[RLIMIT_NPROC];
for (i = 0; i < UCOUNT_COUNTS; i++) {
init_user_ns.ucount_max[i] = max_threads/2;
}
}
int __weak arch_dup_task_struct(struct task_struct *dst,
struct task_struct *src)
{
*dst = *src;
return 0;
}
init/main.c: Give init_task a canary Tasks get their end of stack set to STACK_END_MAGIC with the aim to catch stack overruns. Currently this feature does not apply to init_task. This patch removes this restriction. Note that a similar patch was posted by Prarit Bhargava some time ago but was never merged: http://marc.info/?l=linux-kernel&m=127144305403241&w=2 Signed-off-by: Aaron Tomlin <atomlin@redhat.com> Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org> Acked-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Michael Ellerman <mpe@ellerman.id.au> Cc: aneesh.kumar@linux.vnet.ibm.com Cc: dzickus@redhat.com Cc: bmr@redhat.com Cc: jcastillo@redhat.com Cc: jgh@redhat.com Cc: minchan@kernel.org Cc: tglx@linutronix.de Cc: hannes@cmpxchg.org Cc: Alex Thorlton <athorlton@sgi.com> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Daeseok Youn <daeseok.youn@gmail.com> Cc: David Rientjes <rientjes@google.com> Cc: Fabian Frederick <fabf@skynet.be> Cc: Geert Uytterhoeven <geert@linux-m68k.org> Cc: Jiri Olsa <jolsa@redhat.com> Cc: Kees Cook <keescook@chromium.org> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Masami Hiramatsu <masami.hiramatsu.pt@hitachi.com> Cc: Michael Opdenacker <michael.opdenacker@free-electrons.com> Cc: Paul Mackerras <paulus@samba.org> Cc: Prarit Bhargava <prarit@redhat.com> Cc: Rik van Riel <riel@redhat.com> Cc: Rusty Russell <rusty@rustcorp.com.au> Cc: Seiji Aguchi <seiji.aguchi@hds.com> Cc: Steven Rostedt <rostedt@goodmis.org> Cc: Vladimir Davydov <vdavydov@parallels.com> Cc: Yasuaki Ishimatsu <isimatu.yasuaki@jp.fujitsu.com> Cc: linuxppc-dev@lists.ozlabs.org Link: http://lkml.kernel.org/r/1410527779-8133-2-git-send-email-atomlin@redhat.com Signed-off-by: Ingo Molnar <mingo@kernel.org>
2014-09-12 13:16:17 +00:00
void set_task_stack_end_magic(struct task_struct *tsk)
{
unsigned long *stackend;
stackend = end_of_stack(tsk);
*stackend = STACK_END_MAGIC; /* for overflow detection */
}
static struct task_struct *dup_task_struct(struct task_struct *orig, int node)
{
struct task_struct *tsk;
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
unsigned long *stack;
struct vm_struct *stack_vm_area;
int err;
if (node == NUMA_NO_NODE)
node = tsk_fork_get_node(orig);
tsk = alloc_task_struct_node(node);
if (!tsk)
return NULL;
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
stack = alloc_thread_stack_node(tsk, node);
if (!stack)
fork: fix error handling in dup_task() The function dup_task() may fail at the following function calls in the following order. 0) alloc_task_struct_node() 1) alloc_thread_info_node() 2) arch_dup_task_struct() Error by 0) is not a matter, it can just return. But error by 1) requires releasing task_struct allocated by 0) before it returns. Likewise, error by 2) requires releasing task_struct and thread_info allocated by 0) and 1). The existing error handling calls free_task_struct() and free_thread_info() which do not only release task_struct and thread_info, but also call architecture specific arch_release_task_struct() and arch_release_thread_info(). The problem is that task_struct and thread_info are not fully initialized yet at this point, but arch_release_task_struct() and arch_release_thread_info() are called with them. For example, x86 defines its own arch_release_task_struct() that releases a task_xstate. If alloc_thread_info_node() fails in dup_task(), arch_release_task_struct() is called with task_struct which is just allocated and filled with garbage in this error handling. This actually happened with tools/testing/fault-injection/failcmd.sh # env FAILCMD_TYPE=fail_page_alloc \ ./tools/testing/fault-injection/failcmd.sh --times=100 \ --min-order=0 --ignore-gfp-wait=0 \ -- make -C tools/testing/selftests/ run_tests In order to fix this issue, make free_{task_struct,thread_info}() not to call arch_release_{task_struct,thread_info}() and call arch_release_{task_struct,thread_info}() implicitly where needed. Default arch_release_task_struct() and arch_release_thread_info() are defined as empty by default. So this change only affects the architectures which implement their own arch_release_task_struct() or arch_release_thread_info() as listed below. arch_release_task_struct(): x86, sh arch_release_thread_info(): mn10300, tile Signed-off-by: Akinobu Mita <akinobu.mita@gmail.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: David Howells <dhowells@redhat.com> Cc: Koichi Yasutake <yasutake.koichi@jp.panasonic.com> Cc: Paul Mundt <lethal@linux-sh.org> Cc: Chris Metcalf <cmetcalf@tilera.com> Cc: Salman Qazi <sqazi@google.com> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: Ingo Molnar <mingo@elte.hu> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-07-30 21:42:33 +00:00
goto free_tsk;
stack_vm_area = task_stack_vm_area(tsk);
err = arch_dup_task_struct(tsk, orig);
/*
* arch_dup_task_struct() clobbers the stack-related fields. Make
* sure they're properly initialized before using any stack-related
* functions again.
*/
tsk->stack = stack;
#ifdef CONFIG_VMAP_STACK
tsk->stack_vm_area = stack_vm_area;
#endif
#ifdef CONFIG_THREAD_INFO_IN_TASK
atomic_set(&tsk->stack_refcount, 1);
#endif
if (err)
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
goto free_stack;
#ifdef CONFIG_SECCOMP
/*
* We must handle setting up seccomp filters once we're under
* the sighand lock in case orig has changed between now and
* then. Until then, filter must be NULL to avoid messing up
* the usage counts on the error path calling free_task.
*/
tsk->seccomp.filter = NULL;
#endif
setup_thread_stack(tsk, orig);
clear_user_return_notifier(tsk);
clear_tsk_need_resched(tsk);
init/main.c: Give init_task a canary Tasks get their end of stack set to STACK_END_MAGIC with the aim to catch stack overruns. Currently this feature does not apply to init_task. This patch removes this restriction. Note that a similar patch was posted by Prarit Bhargava some time ago but was never merged: http://marc.info/?l=linux-kernel&m=127144305403241&w=2 Signed-off-by: Aaron Tomlin <atomlin@redhat.com> Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org> Acked-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Michael Ellerman <mpe@ellerman.id.au> Cc: aneesh.kumar@linux.vnet.ibm.com Cc: dzickus@redhat.com Cc: bmr@redhat.com Cc: jcastillo@redhat.com Cc: jgh@redhat.com Cc: minchan@kernel.org Cc: tglx@linutronix.de Cc: hannes@cmpxchg.org Cc: Alex Thorlton <athorlton@sgi.com> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Daeseok Youn <daeseok.youn@gmail.com> Cc: David Rientjes <rientjes@google.com> Cc: Fabian Frederick <fabf@skynet.be> Cc: Geert Uytterhoeven <geert@linux-m68k.org> Cc: Jiri Olsa <jolsa@redhat.com> Cc: Kees Cook <keescook@chromium.org> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Masami Hiramatsu <masami.hiramatsu.pt@hitachi.com> Cc: Michael Opdenacker <michael.opdenacker@free-electrons.com> Cc: Paul Mackerras <paulus@samba.org> Cc: Prarit Bhargava <prarit@redhat.com> Cc: Rik van Riel <riel@redhat.com> Cc: Rusty Russell <rusty@rustcorp.com.au> Cc: Seiji Aguchi <seiji.aguchi@hds.com> Cc: Steven Rostedt <rostedt@goodmis.org> Cc: Vladimir Davydov <vdavydov@parallels.com> Cc: Yasuaki Ishimatsu <isimatu.yasuaki@jp.fujitsu.com> Cc: linuxppc-dev@lists.ozlabs.org Link: http://lkml.kernel.org/r/1410527779-8133-2-git-send-email-atomlin@redhat.com Signed-off-by: Ingo Molnar <mingo@kernel.org>
2014-09-12 13:16:17 +00:00
set_task_stack_end_magic(tsk);
#ifdef CONFIG_CC_STACKPROTECTOR
tsk->stack_canary = get_random_int();
#endif
/*
* One for us, one for whoever does the "release_task()" (usually
* parent)
*/
atomic_set(&tsk->usage, 2);
#ifdef CONFIG_BLK_DEV_IO_TRACE
tsk->btrace_seq = 0;
#endif
tsk->splice_pipe = NULL;
net: use a per task frag allocator We currently use a per socket order-0 page cache for tcp_sendmsg() operations. This page is used to build fragments for skbs. Its done to increase probability of coalescing small write() into single segments in skbs still in write queue (not yet sent) But it wastes a lot of memory for applications handling many mostly idle sockets, since each socket holds one page in sk->sk_sndmsg_page Its also quite inefficient to build TSO 64KB packets, because we need about 16 pages per skb on arches where PAGE_SIZE = 4096, so we hit page allocator more than wanted. This patch adds a per task frag allocator and uses bigger pages, if available. An automatic fallback is done in case of memory pressure. (up to 32768 bytes per frag, thats order-3 pages on x86) This increases TCP stream performance by 20% on loopback device, but also benefits on other network devices, since 8x less frags are mapped on transmit and unmapped on tx completion. Alexander Duyck mentioned a probable performance win on systems with IOMMU enabled. Its possible some SG enabled hardware cant cope with bigger fragments, but their ndo_start_xmit() should already handle this, splitting a fragment in sub fragments, since some arches have PAGE_SIZE=65536 Successfully tested on various ethernet devices. (ixgbe, igb, bnx2x, tg3, mellanox mlx4) Signed-off-by: Eric Dumazet <edumazet@google.com> Cc: Ben Hutchings <bhutchings@solarflare.com> Cc: Vijay Subramanian <subramanian.vijay@gmail.com> Cc: Alexander Duyck <alexander.h.duyck@intel.com> Tested-by: Vijay Subramanian <subramanian.vijay@gmail.com> Signed-off-by: David S. Miller <davem@davemloft.net>
2012-09-23 23:04:42 +00:00
tsk->task_frag.page = NULL;
tsk->wake_q.next = NULL;
account_kernel_stack(tsk, 1);
kernel: add kcov code coverage kcov provides code coverage collection for coverage-guided fuzzing (randomized testing). Coverage-guided fuzzing is a testing technique that uses coverage feedback to determine new interesting inputs to a system. A notable user-space example is AFL (http://lcamtuf.coredump.cx/afl/). However, this technique is not widely used for kernel testing due to missing compiler and kernel support. kcov does not aim to collect as much coverage as possible. It aims to collect more or less stable coverage that is function of syscall inputs. To achieve this goal it does not collect coverage in soft/hard interrupts and instrumentation of some inherently non-deterministic or non-interesting parts of kernel is disbled (e.g. scheduler, locking). Currently there is a single coverage collection mode (tracing), but the API anticipates additional collection modes. Initially I also implemented a second mode which exposes coverage in a fixed-size hash table of counters (what Quentin used in his original patch). I've dropped the second mode for simplicity. This patch adds the necessary support on kernel side. The complimentary compiler support was added in gcc revision 231296. We've used this support to build syzkaller system call fuzzer, which has found 90 kernel bugs in just 2 months: https://github.com/google/syzkaller/wiki/Found-Bugs We've also found 30+ bugs in our internal systems with syzkaller. Another (yet unexplored) direction where kcov coverage would greatly help is more traditional "blob mutation". For example, mounting a random blob as a filesystem, or receiving a random blob over wire. Why not gcov. Typical fuzzing loop looks as follows: (1) reset coverage, (2) execute a bit of code, (3) collect coverage, repeat. A typical coverage can be just a dozen of basic blocks (e.g. an invalid input). In such context gcov becomes prohibitively expensive as reset/collect coverage steps depend on total number of basic blocks/edges in program (in case of kernel it is about 2M). Cost of kcov depends only on number of executed basic blocks/edges. On top of that, kernel requires per-thread coverage because there are always background threads and unrelated processes that also produce coverage. With inlined gcov instrumentation per-thread coverage is not possible. kcov exposes kernel PCs and control flow to user-space which is insecure. But debugfs should not be mapped as user accessible. Based on a patch by Quentin Casasnovas. [akpm@linux-foundation.org: make task_struct.kcov_mode have type `enum kcov_mode'] [akpm@linux-foundation.org: unbreak allmodconfig] [akpm@linux-foundation.org: follow x86 Makefile layout standards] Signed-off-by: Dmitry Vyukov <dvyukov@google.com> Reviewed-by: Kees Cook <keescook@chromium.org> Cc: syzkaller <syzkaller@googlegroups.com> Cc: Vegard Nossum <vegard.nossum@oracle.com> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Tavis Ormandy <taviso@google.com> Cc: Will Deacon <will.deacon@arm.com> Cc: Quentin Casasnovas <quentin.casasnovas@oracle.com> Cc: Kostya Serebryany <kcc@google.com> Cc: Eric Dumazet <edumazet@google.com> Cc: Alexander Potapenko <glider@google.com> Cc: Kees Cook <keescook@google.com> Cc: Bjorn Helgaas <bhelgaas@google.com> Cc: Sasha Levin <sasha.levin@oracle.com> Cc: David Drysdale <drysdale@google.com> Cc: Ard Biesheuvel <ard.biesheuvel@linaro.org> Cc: Andrey Ryabinin <ryabinin.a.a@gmail.com> Cc: Kirill A. Shutemov <kirill@shutemov.name> Cc: Jiri Slaby <jslaby@suse.cz> Cc: Ingo Molnar <mingo@elte.hu> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: "H. Peter Anvin" <hpa@zytor.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-03-22 21:27:30 +00:00
kcov_task_init(tsk);
return tsk;
Clarify naming of thread info/stack allocators We've had the thread info allocated together with the thread stack for most architectures for a long time (since the thread_info was split off from the task struct), but that is about to change. But the patches that move the thread info to be off-stack (and a part of the task struct instead) made it clear how confused the allocator and freeing functions are. Because the common case was that we share an allocation with the thread stack and the thread_info, the two pointers were identical. That identity then meant that we would have things like ti = alloc_thread_info_node(tsk, node); ... tsk->stack = ti; which certainly _worked_ (since stack and thread_info have the same value), but is rather confusing: why are we assigning a thread_info to the stack? And if we move the thread_info away, the "confusing" code just gets to be entirely bogus. So remove all this confusion, and make it clear that we are doing the stack allocation by renaming and clarifying the function names to be about the stack. The fact that the thread_info then shares the allocation is an implementation detail, and not really about the allocation itself. This is a pure renaming and type fix: we pass in the same pointer, it's just that we clarify what the pointer means. The ia64 code that actually only has one single allocation (for all of task_struct, thread_info and kernel thread stack) now looks a bit odd, but since "tsk->stack" is actually not even used there, that oddity doesn't matter. It would be a separate thing to clean that up, I intentionally left the ia64 changes as a pure brute-force renaming and type change. Acked-by: Andy Lutomirski <luto@amacapital.net> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-06-24 22:09:37 +00:00
free_stack:
free_thread_stack(tsk);
fork: fix error handling in dup_task() The function dup_task() may fail at the following function calls in the following order. 0) alloc_task_struct_node() 1) alloc_thread_info_node() 2) arch_dup_task_struct() Error by 0) is not a matter, it can just return. But error by 1) requires releasing task_struct allocated by 0) before it returns. Likewise, error by 2) requires releasing task_struct and thread_info allocated by 0) and 1). The existing error handling calls free_task_struct() and free_thread_info() which do not only release task_struct and thread_info, but also call architecture specific arch_release_task_struct() and arch_release_thread_info(). The problem is that task_struct and thread_info are not fully initialized yet at this point, but arch_release_task_struct() and arch_release_thread_info() are called with them. For example, x86 defines its own arch_release_task_struct() that releases a task_xstate. If alloc_thread_info_node() fails in dup_task(), arch_release_task_struct() is called with task_struct which is just allocated and filled with garbage in this error handling. This actually happened with tools/testing/fault-injection/failcmd.sh # env FAILCMD_TYPE=fail_page_alloc \ ./tools/testing/fault-injection/failcmd.sh --times=100 \ --min-order=0 --ignore-gfp-wait=0 \ -- make -C tools/testing/selftests/ run_tests In order to fix this issue, make free_{task_struct,thread_info}() not to call arch_release_{task_struct,thread_info}() and call arch_release_{task_struct,thread_info}() implicitly where needed. Default arch_release_task_struct() and arch_release_thread_info() are defined as empty by default. So this change only affects the architectures which implement their own arch_release_task_struct() or arch_release_thread_info() as listed below. arch_release_task_struct(): x86, sh arch_release_thread_info(): mn10300, tile Signed-off-by: Akinobu Mita <akinobu.mita@gmail.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: David Howells <dhowells@redhat.com> Cc: Koichi Yasutake <yasutake.koichi@jp.panasonic.com> Cc: Paul Mundt <lethal@linux-sh.org> Cc: Chris Metcalf <cmetcalf@tilera.com> Cc: Salman Qazi <sqazi@google.com> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: Ingo Molnar <mingo@elte.hu> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-07-30 21:42:33 +00:00
free_tsk:
free_task_struct(tsk);
return NULL;
}
#ifdef CONFIG_MMU
static __latent_entropy int dup_mmap(struct mm_struct *mm,
struct mm_struct *oldmm)
{
struct vm_area_struct *mpnt, *tmp, *prev, **pprev;
struct rb_node **rb_link, *rb_parent;
int retval;
unsigned long charge;
2017-02-22 23:42:27 +00:00
LIST_HEAD(uf);
uprobe_start_dup_mmap();
if (down_write_killable(&oldmm->mmap_sem)) {
retval = -EINTR;
goto fail_uprobe_end;
}
flush_cache_dup_mm(oldmm);
uprobe_dup_mmap(oldmm, mm);
/*
* Not linked in yet - no deadlock potential:
*/
down_write_nested(&mm->mmap_sem, SINGLE_DEPTH_NESTING);
/* No ordering required: file already has been exposed. */
RCU_INIT_POINTER(mm->exe_file, get_mm_exe_file(oldmm));
mm->total_vm = oldmm->total_vm;
mm: rework virtual memory accounting When inspecting a vague code inside prctl(PR_SET_MM_MEM) call (which testing the RLIMIT_DATA value to figure out if we're allowed to assign new @start_brk, @brk, @start_data, @end_data from mm_struct) it's been commited that RLIMIT_DATA in a form it's implemented now doesn't do anything useful because most of user-space libraries use mmap() syscall for dynamic memory allocations. Linus suggested to convert RLIMIT_DATA rlimit into something suitable for anonymous memory accounting. But in this patch we go further, and the changes are bundled together as: * keep vma counting if CONFIG_PROC_FS=n, will be used for limits * replace mm->shared_vm with better defined mm->data_vm * account anonymous executable areas as executable * account file-backed growsdown/up areas as stack * drop struct file* argument from vm_stat_account * enforce RLIMIT_DATA for size of data areas This way code looks cleaner: now code/stack/data classification depends only on vm_flags state: VM_EXEC & ~VM_WRITE -> code (VmExe + VmLib in proc) VM_GROWSUP | VM_GROWSDOWN -> stack (VmStk) VM_WRITE & ~VM_SHARED & !stack -> data (VmData) The rest (VmSize - VmData - VmStk - VmExe - VmLib) could be called "shared", but that might be strange beast like readonly-private or VM_IO area. - RLIMIT_AS limits whole address space "VmSize" - RLIMIT_STACK limits stack "VmStk" (but each vma individually) - RLIMIT_DATA now limits "VmData" Signed-off-by: Konstantin Khlebnikov <koct9i@gmail.com> Signed-off-by: Cyrill Gorcunov <gorcunov@openvz.org> Cc: Quentin Casasnovas <quentin.casasnovas@oracle.com> Cc: Vegard Nossum <vegard.nossum@oracle.com> Acked-by: Linus Torvalds <torvalds@linux-foundation.org> Cc: Willy Tarreau <w@1wt.eu> Cc: Andy Lutomirski <luto@amacapital.net> Cc: Kees Cook <keescook@google.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Cc: Pavel Emelyanov <xemul@virtuozzo.com> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-01-14 23:22:07 +00:00
mm->data_vm = oldmm->data_vm;
mm->exec_vm = oldmm->exec_vm;
mm->stack_vm = oldmm->stack_vm;
rb_link = &mm->mm_rb.rb_node;
rb_parent = NULL;
pprev = &mm->mmap;
ksm: the mm interface to ksm This patch presents the mm interface to a dummy version of ksm.c, for better scrutiny of that interface: the real ksm.c follows later. When CONFIG_KSM is not set, madvise(2) reject MADV_MERGEABLE and MADV_UNMERGEABLE with EINVAL, since that seems more helpful than pretending that they can be serviced. But when CONFIG_KSM=y, accept them even if KSM is not currently running, and even on areas which KSM will not touch (e.g. hugetlb or shared file or special driver mappings). Like other madvices, report ENOMEM despite success if any area in the range is unmapped, and use EAGAIN to report out of memory. Define vma flag VM_MERGEABLE to identify an area on which KSM may try merging pages: leave it to ksm_madvise() to decide whether to set it. Define mm flag MMF_VM_MERGEABLE to identify an mm which might contain VM_MERGEABLE areas, to minimize callouts when forking or exiting. Based upon earlier patches by Chris Wright and Izik Eidus. Signed-off-by: Hugh Dickins <hugh.dickins@tiscali.co.uk> Signed-off-by: Chris Wright <chrisw@redhat.com> Signed-off-by: Izik Eidus <ieidus@redhat.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Rik van Riel <riel@redhat.com> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Lee Schermerhorn <lee.schermerhorn@hp.com> Cc: Avi Kivity <avi@redhat.com> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-22 00:01:57 +00:00
retval = ksm_fork(mm, oldmm);
2011-01-13 23:46:58 +00:00
if (retval)
goto out;
retval = khugepaged_fork(mm, oldmm);
ksm: the mm interface to ksm This patch presents the mm interface to a dummy version of ksm.c, for better scrutiny of that interface: the real ksm.c follows later. When CONFIG_KSM is not set, madvise(2) reject MADV_MERGEABLE and MADV_UNMERGEABLE with EINVAL, since that seems more helpful than pretending that they can be serviced. But when CONFIG_KSM=y, accept them even if KSM is not currently running, and even on areas which KSM will not touch (e.g. hugetlb or shared file or special driver mappings). Like other madvices, report ENOMEM despite success if any area in the range is unmapped, and use EAGAIN to report out of memory. Define vma flag VM_MERGEABLE to identify an area on which KSM may try merging pages: leave it to ksm_madvise() to decide whether to set it. Define mm flag MMF_VM_MERGEABLE to identify an mm which might contain VM_MERGEABLE areas, to minimize callouts when forking or exiting. Based upon earlier patches by Chris Wright and Izik Eidus. Signed-off-by: Hugh Dickins <hugh.dickins@tiscali.co.uk> Signed-off-by: Chris Wright <chrisw@redhat.com> Signed-off-by: Izik Eidus <ieidus@redhat.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Rik van Riel <riel@redhat.com> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Lee Schermerhorn <lee.schermerhorn@hp.com> Cc: Avi Kivity <avi@redhat.com> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-22 00:01:57 +00:00
if (retval)
goto out;
prev = NULL;
for (mpnt = oldmm->mmap; mpnt; mpnt = mpnt->vm_next) {
struct file *file;
if (mpnt->vm_flags & VM_DONTCOPY) {
mm: rework virtual memory accounting When inspecting a vague code inside prctl(PR_SET_MM_MEM) call (which testing the RLIMIT_DATA value to figure out if we're allowed to assign new @start_brk, @brk, @start_data, @end_data from mm_struct) it's been commited that RLIMIT_DATA in a form it's implemented now doesn't do anything useful because most of user-space libraries use mmap() syscall for dynamic memory allocations. Linus suggested to convert RLIMIT_DATA rlimit into something suitable for anonymous memory accounting. But in this patch we go further, and the changes are bundled together as: * keep vma counting if CONFIG_PROC_FS=n, will be used for limits * replace mm->shared_vm with better defined mm->data_vm * account anonymous executable areas as executable * account file-backed growsdown/up areas as stack * drop struct file* argument from vm_stat_account * enforce RLIMIT_DATA for size of data areas This way code looks cleaner: now code/stack/data classification depends only on vm_flags state: VM_EXEC & ~VM_WRITE -> code (VmExe + VmLib in proc) VM_GROWSUP | VM_GROWSDOWN -> stack (VmStk) VM_WRITE & ~VM_SHARED & !stack -> data (VmData) The rest (VmSize - VmData - VmStk - VmExe - VmLib) could be called "shared", but that might be strange beast like readonly-private or VM_IO area. - RLIMIT_AS limits whole address space "VmSize" - RLIMIT_STACK limits stack "VmStk" (but each vma individually) - RLIMIT_DATA now limits "VmData" Signed-off-by: Konstantin Khlebnikov <koct9i@gmail.com> Signed-off-by: Cyrill Gorcunov <gorcunov@openvz.org> Cc: Quentin Casasnovas <quentin.casasnovas@oracle.com> Cc: Vegard Nossum <vegard.nossum@oracle.com> Acked-by: Linus Torvalds <torvalds@linux-foundation.org> Cc: Willy Tarreau <w@1wt.eu> Cc: Andy Lutomirski <luto@amacapital.net> Cc: Kees Cook <keescook@google.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Cc: Pavel Emelyanov <xemul@virtuozzo.com> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-01-14 23:22:07 +00:00
vm_stat_account(mm, mpnt->vm_flags, -vma_pages(mpnt));
continue;
}
charge = 0;
if (mpnt->vm_flags & VM_ACCOUNT) {
unsigned long len = vma_pages(mpnt);
if (security_vm_enough_memory_mm(oldmm, len)) /* sic */
goto fail_nomem;
charge = len;
}
tmp = kmem_cache_alloc(vm_area_cachep, GFP_KERNEL);
if (!tmp)
goto fail_nomem;
*tmp = *mpnt;
mm: change anon_vma linking to fix multi-process server scalability issue The old anon_vma code can lead to scalability issues with heavily forking workloads. Specifically, each anon_vma will be shared between the parent process and all its child processes. In a workload with 1000 child processes and a VMA with 1000 anonymous pages per process that get COWed, this leads to a system with a million anonymous pages in the same anon_vma, each of which is mapped in just one of the 1000 processes. However, the current rmap code needs to walk them all, leading to O(N) scanning complexity for each page. This can result in systems where one CPU is walking the page tables of 1000 processes in page_referenced_one, while all other CPUs are stuck on the anon_vma lock. This leads to catastrophic failure for a benchmark like AIM7, where the total number of processes can reach in the tens of thousands. Real workloads are still a factor 10 less process intensive than AIM7, but they are catching up. This patch changes the way anon_vmas and VMAs are linked, which allows us to associate multiple anon_vmas with a VMA. At fork time, each child process gets its own anon_vmas, in which its COWed pages will be instantiated. The parents' anon_vma is also linked to the VMA, because non-COWed pages could be present in any of the children. This reduces rmap scanning complexity to O(1) for the pages of the 1000 child processes, with O(N) complexity for at most 1/N pages in the system. This reduces the average scanning cost in heavily forking workloads from O(N) to 2. The only real complexity in this patch stems from the fact that linking a VMA to anon_vmas now involves memory allocations. This means vma_adjust can fail, if it needs to attach a VMA to anon_vma structures. This in turn means error handling needs to be added to the calling functions. A second source of complexity is that, because there can be multiple anon_vmas, the anon_vma linking in vma_adjust can no longer be done under "the" anon_vma lock. To prevent the rmap code from walking up an incomplete VMA, this patch introduces the VM_LOCK_RMAP VMA flag. This bit flag uses the same slot as the NOMMU VM_MAPPED_COPY, with an ifdef in mm.h to make sure it is impossible to compile a kernel that needs both symbolic values for the same bitflag. Some test results: Without the anon_vma changes, when AIM7 hits around 9.7k users (on a test box with 16GB RAM and not quite enough IO), the system ends up running >99% in system time, with every CPU on the same anon_vma lock in the pageout code. With these changes, AIM7 hits the cross-over point around 29.7k users. This happens with ~99% IO wait time, there never seems to be any spike in system time. The anon_vma lock contention appears to be resolved. [akpm@linux-foundation.org: cleanups] Signed-off-by: Rik van Riel <riel@redhat.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Larry Woodman <lwoodman@redhat.com> Cc: Lee Schermerhorn <Lee.Schermerhorn@hp.com> Cc: Minchan Kim <minchan.kim@gmail.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-03-05 21:42:07 +00:00
INIT_LIST_HEAD(&tmp->anon_vma_chain);
retval = vma_dup_policy(mpnt, tmp);
if (retval)
goto fail_nomem_policy;
tmp->vm_mm = mm;
2017-02-22 23:42:27 +00:00
retval = dup_userfaultfd(tmp, &uf);
if (retval)
goto fail_nomem_anon_vma_fork;
mm: change anon_vma linking to fix multi-process server scalability issue The old anon_vma code can lead to scalability issues with heavily forking workloads. Specifically, each anon_vma will be shared between the parent process and all its child processes. In a workload with 1000 child processes and a VMA with 1000 anonymous pages per process that get COWed, this leads to a system with a million anonymous pages in the same anon_vma, each of which is mapped in just one of the 1000 processes. However, the current rmap code needs to walk them all, leading to O(N) scanning complexity for each page. This can result in systems where one CPU is walking the page tables of 1000 processes in page_referenced_one, while all other CPUs are stuck on the anon_vma lock. This leads to catastrophic failure for a benchmark like AIM7, where the total number of processes can reach in the tens of thousands. Real workloads are still a factor 10 less process intensive than AIM7, but they are catching up. This patch changes the way anon_vmas and VMAs are linked, which allows us to associate multiple anon_vmas with a VMA. At fork time, each child process gets its own anon_vmas, in which its COWed pages will be instantiated. The parents' anon_vma is also linked to the VMA, because non-COWed pages could be present in any of the children. This reduces rmap scanning complexity to O(1) for the pages of the 1000 child processes, with O(N) complexity for at most 1/N pages in the system. This reduces the average scanning cost in heavily forking workloads from O(N) to 2. The only real complexity in this patch stems from the fact that linking a VMA to anon_vmas now involves memory allocations. This means vma_adjust can fail, if it needs to attach a VMA to anon_vma structures. This in turn means error handling needs to be added to the calling functions. A second source of complexity is that, because there can be multiple anon_vmas, the anon_vma linking in vma_adjust can no longer be done under "the" anon_vma lock. To prevent the rmap code from walking up an incomplete VMA, this patch introduces the VM_LOCK_RMAP VMA flag. This bit flag uses the same slot as the NOMMU VM_MAPPED_COPY, with an ifdef in mm.h to make sure it is impossible to compile a kernel that needs both symbolic values for the same bitflag. Some test results: Without the anon_vma changes, when AIM7 hits around 9.7k users (on a test box with 16GB RAM and not quite enough IO), the system ends up running >99% in system time, with every CPU on the same anon_vma lock in the pageout code. With these changes, AIM7 hits the cross-over point around 29.7k users. This happens with ~99% IO wait time, there never seems to be any spike in system time. The anon_vma lock contention appears to be resolved. [akpm@linux-foundation.org: cleanups] Signed-off-by: Rik van Riel <riel@redhat.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Larry Woodman <lwoodman@redhat.com> Cc: Lee Schermerhorn <Lee.Schermerhorn@hp.com> Cc: Minchan Kim <minchan.kim@gmail.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-03-05 21:42:07 +00:00
if (anon_vma_fork(tmp, mpnt))
goto fail_nomem_anon_vma_fork;
2017-02-22 23:42:27 +00:00
tmp->vm_flags &= ~(VM_LOCKED | VM_LOCKONFAULT);
tmp->vm_next = tmp->vm_prev = NULL;
file = tmp->vm_file;
if (file) {
struct inode *inode = file_inode(file);
struct address_space *mapping = file->f_mapping;
get_file(file);
if (tmp->vm_flags & VM_DENYWRITE)
atomic_dec(&inode->i_writecount);
i_mmap_lock_write(mapping);
if (tmp->vm_flags & VM_SHARED)
atomic_inc(&mapping->i_mmap_writable);
flush_dcache_mmap_lock(mapping);
/* insert tmp into the share list, just after mpnt */
vma_interval_tree_insert_after(tmp, mpnt,
&mapping->i_mmap);
flush_dcache_mmap_unlock(mapping);
i_mmap_unlock_write(mapping);
}
hugetlb: reserve huge pages for reliable MAP_PRIVATE hugetlbfs mappings until fork() This patch reserves huge pages at mmap() time for MAP_PRIVATE mappings in a similar manner to the reservations taken for MAP_SHARED mappings. The reserve count is accounted both globally and on a per-VMA basis for private mappings. This guarantees that a process that successfully calls mmap() will successfully fault all pages in the future unless fork() is called. The characteristics of private mappings of hugetlbfs files behaviour after this patch are; 1. The process calling mmap() is guaranteed to succeed all future faults until it forks(). 2. On fork(), the parent may die due to SIGKILL on writes to the private mapping if enough pages are not available for the COW. For reasonably reliable behaviour in the face of a small huge page pool, children of hugepage-aware processes should not reference the mappings; such as might occur when fork()ing to exec(). 3. On fork(), the child VMAs inherit no reserves. Reads on pages already faulted by the parent will succeed. Successful writes will depend on enough huge pages being free in the pool. 4. Quotas of the hugetlbfs mount are checked at reserve time for the mapper and at fault time otherwise. Before this patch, all reads or writes in the child potentially needs page allocations that can later lead to the death of the parent. This applies to reads and writes of uninstantiated pages as well as COW. After the patch it is only a write to an instantiated page that causes problems. Signed-off-by: Mel Gorman <mel@csn.ul.ie> Acked-by: Adam Litke <agl@us.ibm.com> Cc: Andy Whitcroft <apw@shadowen.org> Cc: William Lee Irwin III <wli@holomorphy.com> Cc: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-07-24 04:27:23 +00:00
/*
* Clear hugetlb-related page reserves for children. This only
* affects MAP_PRIVATE mappings. Faults generated by the child
* are not guaranteed to succeed, even if read-only
*/
if (is_vm_hugetlb_page(tmp))
reset_vma_resv_huge_pages(tmp);
/*
* Link in the new vma and copy the page table entries.
*/
*pprev = tmp;
pprev = &tmp->vm_next;
tmp->vm_prev = prev;
prev = tmp;
__vma_link_rb(mm, tmp, rb_link, rb_parent);
rb_link = &tmp->vm_rb.rb_right;
rb_parent = &tmp->vm_rb;
mm->map_count++;
retval = copy_page_range(mm, oldmm, mpnt);
if (tmp->vm_ops && tmp->vm_ops->open)
tmp->vm_ops->open(tmp);
if (retval)
goto out;
}
/* a new mm has just been created */
arch_dup_mmap(oldmm, mm);
retval = 0;
out:
up_write(&mm->mmap_sem);
flush_tlb_mm(oldmm);
up_write(&oldmm->mmap_sem);
2017-02-22 23:42:27 +00:00
dup_userfaultfd_complete(&uf);
fail_uprobe_end:
uprobe_end_dup_mmap();
return retval;
mm: change anon_vma linking to fix multi-process server scalability issue The old anon_vma code can lead to scalability issues with heavily forking workloads. Specifically, each anon_vma will be shared between the parent process and all its child processes. In a workload with 1000 child processes and a VMA with 1000 anonymous pages per process that get COWed, this leads to a system with a million anonymous pages in the same anon_vma, each of which is mapped in just one of the 1000 processes. However, the current rmap code needs to walk them all, leading to O(N) scanning complexity for each page. This can result in systems where one CPU is walking the page tables of 1000 processes in page_referenced_one, while all other CPUs are stuck on the anon_vma lock. This leads to catastrophic failure for a benchmark like AIM7, where the total number of processes can reach in the tens of thousands. Real workloads are still a factor 10 less process intensive than AIM7, but they are catching up. This patch changes the way anon_vmas and VMAs are linked, which allows us to associate multiple anon_vmas with a VMA. At fork time, each child process gets its own anon_vmas, in which its COWed pages will be instantiated. The parents' anon_vma is also linked to the VMA, because non-COWed pages could be present in any of the children. This reduces rmap scanning complexity to O(1) for the pages of the 1000 child processes, with O(N) complexity for at most 1/N pages in the system. This reduces the average scanning cost in heavily forking workloads from O(N) to 2. The only real complexity in this patch stems from the fact that linking a VMA to anon_vmas now involves memory allocations. This means vma_adjust can fail, if it needs to attach a VMA to anon_vma structures. This in turn means error handling needs to be added to the calling functions. A second source of complexity is that, because there can be multiple anon_vmas, the anon_vma linking in vma_adjust can no longer be done under "the" anon_vma lock. To prevent the rmap code from walking up an incomplete VMA, this patch introduces the VM_LOCK_RMAP VMA flag. This bit flag uses the same slot as the NOMMU VM_MAPPED_COPY, with an ifdef in mm.h to make sure it is impossible to compile a kernel that needs both symbolic values for the same bitflag. Some test results: Without the anon_vma changes, when AIM7 hits around 9.7k users (on a test box with 16GB RAM and not quite enough IO), the system ends up running >99% in system time, with every CPU on the same anon_vma lock in the pageout code. With these changes, AIM7 hits the cross-over point around 29.7k users. This happens with ~99% IO wait time, there never seems to be any spike in system time. The anon_vma lock contention appears to be resolved. [akpm@linux-foundation.org: cleanups] Signed-off-by: Rik van Riel <riel@redhat.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Larry Woodman <lwoodman@redhat.com> Cc: Lee Schermerhorn <Lee.Schermerhorn@hp.com> Cc: Minchan Kim <minchan.kim@gmail.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-03-05 21:42:07 +00:00
fail_nomem_anon_vma_fork:
mpol_put(vma_policy(tmp));
fail_nomem_policy:
kmem_cache_free(vm_area_cachep, tmp);
fail_nomem:
retval = -ENOMEM;
vm_unacct_memory(charge);
goto out;
}
static inline int mm_alloc_pgd(struct mm_struct *mm)
{
mm->pgd = pgd_alloc(mm);
if (unlikely(!mm->pgd))
return -ENOMEM;
return 0;
}
static inline void mm_free_pgd(struct mm_struct *mm)
{
pgd_free(mm, mm->pgd);
}
#else
static int dup_mmap(struct mm_struct *mm, struct mm_struct *oldmm)
{
down_write(&oldmm->mmap_sem);
RCU_INIT_POINTER(mm->exe_file, get_mm_exe_file(oldmm));
up_write(&oldmm->mmap_sem);
return 0;
}
#define mm_alloc_pgd(mm) (0)
#define mm_free_pgd(mm)
#endif /* CONFIG_MMU */
__cacheline_aligned_in_smp DEFINE_SPINLOCK(mmlist_lock);
#define allocate_mm() (kmem_cache_alloc(mm_cachep, GFP_KERNEL))
#define free_mm(mm) (kmem_cache_free(mm_cachep, (mm)))
static unsigned long default_dump_filter = MMF_DUMP_FILTER_DEFAULT;
static int __init coredump_filter_setup(char *s)
{
default_dump_filter =
(simple_strtoul(s, NULL, 0) << MMF_DUMP_FILTER_SHIFT) &
MMF_DUMP_FILTER_MASK;
return 1;
}
__setup("coredump_filter=", coredump_filter_setup);
#include <linux/init_task.h>
static void mm_init_aio(struct mm_struct *mm)
{
#ifdef CONFIG_AIO
spin_lock_init(&mm->ioctx_lock);
aio: convert the ioctx list to table lookup v3 On Wed, Jun 12, 2013 at 11:14:40AM -0700, Kent Overstreet wrote: > On Mon, Apr 15, 2013 at 02:40:55PM +0300, Octavian Purdila wrote: > > When using a large number of threads performing AIO operations the > > IOCTX list may get a significant number of entries which will cause > > significant overhead. For example, when running this fio script: > > > > rw=randrw; size=256k ;directory=/mnt/fio; ioengine=libaio; iodepth=1 > > blocksize=1024; numjobs=512; thread; loops=100 > > > > on an EXT2 filesystem mounted on top of a ramdisk we can observe up to > > 30% CPU time spent by lookup_ioctx: > > > > 32.51% [guest.kernel] [g] lookup_ioctx > > 9.19% [guest.kernel] [g] __lock_acquire.isra.28 > > 4.40% [guest.kernel] [g] lock_release > > 4.19% [guest.kernel] [g] sched_clock_local > > 3.86% [guest.kernel] [g] local_clock > > 3.68% [guest.kernel] [g] native_sched_clock > > 3.08% [guest.kernel] [g] sched_clock_cpu > > 2.64% [guest.kernel] [g] lock_release_holdtime.part.11 > > 2.60% [guest.kernel] [g] memcpy > > 2.33% [guest.kernel] [g] lock_acquired > > 2.25% [guest.kernel] [g] lock_acquire > > 1.84% [guest.kernel] [g] do_io_submit > > > > This patchs converts the ioctx list to a radix tree. For a performance > > comparison the above FIO script was run on a 2 sockets 8 core > > machine. This are the results (average and %rsd of 10 runs) for the > > original list based implementation and for the radix tree based > > implementation: > > > > cores 1 2 4 8 16 32 > > list 109376 ms 69119 ms 35682 ms 22671 ms 19724 ms 16408 ms > > %rsd 0.69% 1.15% 1.17% 1.21% 1.71% 1.43% > > radix 73651 ms 41748 ms 23028 ms 16766 ms 15232 ms 13787 ms > > %rsd 1.19% 0.98% 0.69% 1.13% 0.72% 0.75% > > % of radix > > relative 66.12% 65.59% 66.63% 72.31% 77.26% 83.66% > > to list > > > > To consider the impact of the patch on the typical case of having > > only one ctx per process the following FIO script was run: > > > > rw=randrw; size=100m ;directory=/mnt/fio; ioengine=libaio; iodepth=1 > > blocksize=1024; numjobs=1; thread; loops=100 > > > > on the same system and the results are the following: > > > > list 58892 ms > > %rsd 0.91% > > radix 59404 ms > > %rsd 0.81% > > % of radix > > relative 100.87% > > to list > > So, I was just doing some benchmarking/profiling to get ready to send > out the aio patches I've got for 3.11 - and it looks like your patch is > causing a ~1.5% throughput regression in my testing :/ ... <snip> I've got an alternate approach for fixing this wart in lookup_ioctx()... Instead of using an rbtree, just use the reserved id in the ring buffer header to index an array pointing the ioctx. It's not finished yet, and it needs to be tidied up, but is most of the way there. -ben -- "Thought is the essence of where you are now." -- kmo> And, a rework of Ben's code, but this was entirely his idea kmo> -Kent bcrl> And fix the code to use the right mm_struct in kill_ioctx(), actually free memory. Signed-off-by: Benjamin LaHaise <bcrl@kvack.org>
2013-07-30 16:54:40 +00:00
mm->ioctx_table = NULL;
#endif
}
static void mm_init_owner(struct mm_struct *mm, struct task_struct *p)
{
#ifdef CONFIG_MEMCG
mm->owner = p;
#endif
}
mm: Add a user_ns owner to mm_struct and fix ptrace permission checks During exec dumpable is cleared if the file that is being executed is not readable by the user executing the file. A bug in ptrace_may_access allows reading the file if the executable happens to enter into a subordinate user namespace (aka clone(CLONE_NEWUSER), unshare(CLONE_NEWUSER), or setns(fd, CLONE_NEWUSER). This problem is fixed with only necessary userspace breakage by adding a user namespace owner to mm_struct, captured at the time of exec, so it is clear in which user namespace CAP_SYS_PTRACE must be present in to be able to safely give read permission to the executable. The function ptrace_may_access is modified to verify that the ptracer has CAP_SYS_ADMIN in task->mm->user_ns instead of task->cred->user_ns. This ensures that if the task changes it's cred into a subordinate user namespace it does not become ptraceable. The function ptrace_attach is modified to only set PT_PTRACE_CAP when CAP_SYS_PTRACE is held over task->mm->user_ns. The intent of PT_PTRACE_CAP is to be a flag to note that whatever permission changes the task might go through the tracer has sufficient permissions for it not to be an issue. task->cred->user_ns is always the same as or descendent of mm->user_ns. Which guarantees that having CAP_SYS_PTRACE over mm->user_ns is the worst case for the tasks credentials. To prevent regressions mm->dumpable and mm->user_ns are not considered when a task has no mm. As simply failing ptrace_may_attach causes regressions in privileged applications attempting to read things such as /proc/<pid>/stat Cc: stable@vger.kernel.org Acked-by: Kees Cook <keescook@chromium.org> Tested-by: Cyrill Gorcunov <gorcunov@openvz.org> Fixes: 8409cca70561 ("userns: allow ptrace from non-init user namespaces") Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2016-10-14 02:23:16 +00:00
static struct mm_struct *mm_init(struct mm_struct *mm, struct task_struct *p,
struct user_namespace *user_ns)
{
mm->mmap = NULL;
mm->mm_rb = RB_ROOT;
mm->vmacache_seqnum = 0;
atomic_set(&mm->mm_users, 1);
atomic_set(&mm->mm_count, 1);
init_rwsem(&mm->mmap_sem);
INIT_LIST_HEAD(&mm->mmlist);
mm->core_state = NULL;
atomic_long_set(&mm->nr_ptes, 0);
mm_nr_pmds_init(mm);
mm->map_count = 0;
mm->locked_vm = 0;
mm->pinned_vm = 0;
memset(&mm->rss_stat, 0, sizeof(mm->rss_stat));
spin_lock_init(&mm->page_table_lock);
mm_init_cpumask(mm);
mm_init_aio(mm);
cgroups: add an owner to the mm_struct Remove the mem_cgroup member from mm_struct and instead adds an owner. This approach was suggested by Paul Menage. The advantage of this approach is that, once the mm->owner is known, using the subsystem id, the cgroup can be determined. It also allows several control groups that are virtually grouped by mm_struct, to exist independent of the memory controller i.e., without adding mem_cgroup's for each controller, to mm_struct. A new config option CONFIG_MM_OWNER is added and the memory resource controller selects this config option. This patch also adds cgroup callbacks to notify subsystems when mm->owner changes. The mm_cgroup_changed callback is called with the task_lock() of the new task held and is called just prior to changing the mm->owner. I am indebted to Paul Menage for the several reviews of this patchset and helping me make it lighter and simpler. This patch was tested on a powerpc box, it was compiled with both the MM_OWNER config turned on and off. After the thread group leader exits, it's moved to init_css_state by cgroup_exit(), thus all future charges from runnings threads would be redirected to the init_css_set's subsystem. Signed-off-by: Balbir Singh <balbir@linux.vnet.ibm.com> Cc: Pavel Emelianov <xemul@openvz.org> Cc: Hugh Dickins <hugh@veritas.com> Cc: Sudhir Kumar <skumar@linux.vnet.ibm.com> Cc: YAMAMOTO Takashi <yamamoto@valinux.co.jp> Cc: Hirokazu Takahashi <taka@valinux.co.jp> Cc: David Rientjes <rientjes@google.com>, Cc: Balbir Singh <balbir@linux.vnet.ibm.com> Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Acked-by: Pekka Enberg <penberg@cs.helsinki.fi> Reviewed-by: Paul Menage <menage@google.com> Cc: Oleg Nesterov <oleg@tv-sign.ru> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-04-29 08:00:16 +00:00
mm_init_owner(mm, p);
mmu_notifier_mm_init(mm);
mm: fix TLB flush race between migration, and change_protection_range There are a few subtle races, between change_protection_range (used by mprotect and change_prot_numa) on one side, and NUMA page migration and compaction on the other side. The basic race is that there is a time window between when the PTE gets made non-present (PROT_NONE or NUMA), and the TLB is flushed. During that time, a CPU may continue writing to the page. This is fine most of the time, however compaction or the NUMA migration code may come in, and migrate the page away. When that happens, the CPU may continue writing, through the cached translation, to what is no longer the current memory location of the process. This only affects x86, which has a somewhat optimistic pte_accessible. All other architectures appear to be safe, and will either always flush, or flush whenever there is a valid mapping, even with no permissions (SPARC). The basic race looks like this: CPU A CPU B CPU C load TLB entry make entry PTE/PMD_NUMA fault on entry read/write old page start migrating page change PTE/PMD to new page read/write old page [*] flush TLB reload TLB from new entry read/write new page lose data [*] the old page may belong to a new user at this point! The obvious fix is to flush remote TLB entries, by making sure that pte_accessible aware of the fact that PROT_NONE and PROT_NUMA memory may still be accessible if there is a TLB flush pending for the mm. This should fix both NUMA migration and compaction. [mgorman@suse.de: fix build] Signed-off-by: Rik van Riel <riel@redhat.com> Signed-off-by: Mel Gorman <mgorman@suse.de> Cc: Alex Thorlton <athorlton@sgi.com> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-12-19 01:08:44 +00:00
clear_tlb_flush_pending(mm);
#if defined(CONFIG_TRANSPARENT_HUGEPAGE) && !USE_SPLIT_PMD_PTLOCKS
mm->pmd_huge_pte = NULL;
#endif
if (current->mm) {
mm->flags = current->mm->flags & MMF_INIT_MASK;
mm->def_flags = current->mm->def_flags & VM_INIT_DEF_MASK;
} else {
mm->flags = default_dump_filter;
mm->def_flags = 0;
}
if (mm_alloc_pgd(mm))
goto fail_nopgd;
if (init_new_context(p, mm))
goto fail_nocontext;
mm: Add a user_ns owner to mm_struct and fix ptrace permission checks During exec dumpable is cleared if the file that is being executed is not readable by the user executing the file. A bug in ptrace_may_access allows reading the file if the executable happens to enter into a subordinate user namespace (aka clone(CLONE_NEWUSER), unshare(CLONE_NEWUSER), or setns(fd, CLONE_NEWUSER). This problem is fixed with only necessary userspace breakage by adding a user namespace owner to mm_struct, captured at the time of exec, so it is clear in which user namespace CAP_SYS_PTRACE must be present in to be able to safely give read permission to the executable. The function ptrace_may_access is modified to verify that the ptracer has CAP_SYS_ADMIN in task->mm->user_ns instead of task->cred->user_ns. This ensures that if the task changes it's cred into a subordinate user namespace it does not become ptraceable. The function ptrace_attach is modified to only set PT_PTRACE_CAP when CAP_SYS_PTRACE is held over task->mm->user_ns. The intent of PT_PTRACE_CAP is to be a flag to note that whatever permission changes the task might go through the tracer has sufficient permissions for it not to be an issue. task->cred->user_ns is always the same as or descendent of mm->user_ns. Which guarantees that having CAP_SYS_PTRACE over mm->user_ns is the worst case for the tasks credentials. To prevent regressions mm->dumpable and mm->user_ns are not considered when a task has no mm. As simply failing ptrace_may_attach causes regressions in privileged applications attempting to read things such as /proc/<pid>/stat Cc: stable@vger.kernel.org Acked-by: Kees Cook <keescook@chromium.org> Tested-by: Cyrill Gorcunov <gorcunov@openvz.org> Fixes: 8409cca70561 ("userns: allow ptrace from non-init user namespaces") Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2016-10-14 02:23:16 +00:00
mm->user_ns = get_user_ns(user_ns);
return mm;
fail_nocontext:
mm_free_pgd(mm);
fail_nopgd:
free_mm(mm);
return NULL;
}
static void check_mm(struct mm_struct *mm)
{
int i;
for (i = 0; i < NR_MM_COUNTERS; i++) {
long x = atomic_long_read(&mm->rss_stat.count[i]);
if (unlikely(x))
printk(KERN_ALERT "BUG: Bad rss-counter state "
"mm:%p idx:%d val:%ld\n", mm, i, x);
}
mm: fix false-positive warning on exit due mm_nr_pmds(mm) The problem is that we check nr_ptes/nr_pmds in exit_mmap() which happens *before* pgd_free(). And if an arch does pte/pmd allocation in pgd_alloc() and frees them in pgd_free() we see offset in counters by the time of the checks. We tried to workaround this by offsetting expected counter value according to FIRST_USER_ADDRESS for both nr_pte and nr_pmd in exit_mmap(). But it doesn't work in some cases: 1. ARM with LPAE enabled also has non-zero USER_PGTABLES_CEILING, but upper addresses occupied with huge pmd entries, so the trick with offsetting expected counter value will get really ugly: we will have to apply it nr_pmds, but not nr_ptes. 2. Metag has non-zero FIRST_USER_ADDRESS, but doesn't do allocation pte/pmd page tables allocation in pgd_alloc(), just setup a pgd entry which is allocated at boot and shared accross all processes. The proposal is to move the check to check_mm() which happens *after* pgd_free() and do proper accounting during pgd_alloc() and pgd_free() which would bring counters to zero if nothing leaked. Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Reported-by: Tyler Baker <tyler.baker@linaro.org> Tested-by: Tyler Baker <tyler.baker@linaro.org> Tested-by: Nishanth Menon <nm@ti.com> Cc: Russell King <linux@arm.linux.org.uk> Cc: James Hogan <james.hogan@imgtec.com> Cc: Guan Xuetao <gxt@mprc.pku.edu.cn> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-02-11 23:26:53 +00:00
if (atomic_long_read(&mm->nr_ptes))
pr_alert("BUG: non-zero nr_ptes on freeing mm: %ld\n",
atomic_long_read(&mm->nr_ptes));
if (mm_nr_pmds(mm))
pr_alert("BUG: non-zero nr_pmds on freeing mm: %ld\n",
mm_nr_pmds(mm));
mm: implement split page table lock for PMD level The basic idea is the same as with PTE level: the lock is embedded into struct page of table's page. We can't use mm->pmd_huge_pte to store pgtables for THP, since we don't take mm->page_table_lock anymore. Let's reuse page->lru of table's page for that. pgtable_pmd_page_ctor() returns true, if initialization is successful and false otherwise. Current implementation never fails, but assumption that constructor can fail will help to port it to -rt where spinlock_t is rather huge and cannot be embedded into struct page -- dynamic allocation is required. Signed-off-by: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Tested-by: Alex Thorlton <athorlton@sgi.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: "Eric W . Biederman" <ebiederm@xmission.com> Cc: "Paul E . McKenney" <paulmck@linux.vnet.ibm.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Andi Kleen <ak@linux.intel.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: Dave Jones <davej@redhat.com> Cc: David Howells <dhowells@redhat.com> Cc: Frederic Weisbecker <fweisbec@gmail.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Kees Cook <keescook@chromium.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rik van Riel <riel@redhat.com> Cc: Robin Holt <robinmholt@gmail.com> Cc: Sedat Dilek <sedat.dilek@gmail.com> Cc: Srikar Dronamraju <srikar@linux.vnet.ibm.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Hugh Dickins <hughd@google.com> Reviewed-by: Steven Rostedt <rostedt@goodmis.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-11-14 22:31:07 +00:00
#if defined(CONFIG_TRANSPARENT_HUGEPAGE) && !USE_SPLIT_PMD_PTLOCKS
VM_BUG_ON_MM(mm->pmd_huge_pte, mm);
#endif
}
/*
* Allocate and initialize an mm_struct.
*/
struct mm_struct *mm_alloc(void)
{
struct mm_struct *mm;
mm = allocate_mm();
if (!mm)
return NULL;
memset(mm, 0, sizeof(*mm));
mm: Add a user_ns owner to mm_struct and fix ptrace permission checks During exec dumpable is cleared if the file that is being executed is not readable by the user executing the file. A bug in ptrace_may_access allows reading the file if the executable happens to enter into a subordinate user namespace (aka clone(CLONE_NEWUSER), unshare(CLONE_NEWUSER), or setns(fd, CLONE_NEWUSER). This problem is fixed with only necessary userspace breakage by adding a user namespace owner to mm_struct, captured at the time of exec, so it is clear in which user namespace CAP_SYS_PTRACE must be present in to be able to safely give read permission to the executable. The function ptrace_may_access is modified to verify that the ptracer has CAP_SYS_ADMIN in task->mm->user_ns instead of task->cred->user_ns. This ensures that if the task changes it's cred into a subordinate user namespace it does not become ptraceable. The function ptrace_attach is modified to only set PT_PTRACE_CAP when CAP_SYS_PTRACE is held over task->mm->user_ns. The intent of PT_PTRACE_CAP is to be a flag to note that whatever permission changes the task might go through the tracer has sufficient permissions for it not to be an issue. task->cred->user_ns is always the same as or descendent of mm->user_ns. Which guarantees that having CAP_SYS_PTRACE over mm->user_ns is the worst case for the tasks credentials. To prevent regressions mm->dumpable and mm->user_ns are not considered when a task has no mm. As simply failing ptrace_may_attach causes regressions in privileged applications attempting to read things such as /proc/<pid>/stat Cc: stable@vger.kernel.org Acked-by: Kees Cook <keescook@chromium.org> Tested-by: Cyrill Gorcunov <gorcunov@openvz.org> Fixes: 8409cca70561 ("userns: allow ptrace from non-init user namespaces") Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2016-10-14 02:23:16 +00:00
return mm_init(mm, current, current_user_ns());
}
/*
* Called when the last reference to the mm
* is dropped: either by a lazy thread or by
* mmput. Free the page directory and the mm.
*/
void __mmdrop(struct mm_struct *mm)
{
BUG_ON(mm == &init_mm);
mm_free_pgd(mm);
destroy_context(mm);
mmu-notifiers: core With KVM/GFP/XPMEM there isn't just the primary CPU MMU pointing to pages. There are secondary MMUs (with secondary sptes and secondary tlbs) too. sptes in the kvm case are shadow pagetables, but when I say spte in mmu-notifier context, I mean "secondary pte". In GRU case there's no actual secondary pte and there's only a secondary tlb because the GRU secondary MMU has no knowledge about sptes and every secondary tlb miss event in the MMU always generates a page fault that has to be resolved by the CPU (this is not the case of KVM where the a secondary tlb miss will walk sptes in hardware and it will refill the secondary tlb transparently to software if the corresponding spte is present). The same way zap_page_range has to invalidate the pte before freeing the page, the spte (and secondary tlb) must also be invalidated before any page is freed and reused. Currently we take a page_count pin on every page mapped by sptes, but that means the pages can't be swapped whenever they're mapped by any spte because they're part of the guest working set. Furthermore a spte unmap event can immediately lead to a page to be freed when the pin is released (so requiring the same complex and relatively slow tlb_gather smp safe logic we have in zap_page_range and that can be avoided completely if the spte unmap event doesn't require an unpin of the page previously mapped in the secondary MMU). The mmu notifiers allow kvm/GRU/XPMEM to attach to the tsk->mm and know when the VM is swapping or freeing or doing anything on the primary MMU so that the secondary MMU code can drop sptes before the pages are freed, avoiding all page pinning and allowing 100% reliable swapping of guest physical address space. Furthermore it avoids the code that teardown the mappings of the secondary MMU, to implement a logic like tlb_gather in zap_page_range that would require many IPI to flush other cpu tlbs, for each fixed number of spte unmapped. To make an example: if what happens on the primary MMU is a protection downgrade (from writeable to wrprotect) the secondary MMU mappings will be invalidated, and the next secondary-mmu-page-fault will call get_user_pages and trigger a do_wp_page through get_user_pages if it called get_user_pages with write=1, and it'll re-establishing an updated spte or secondary-tlb-mapping on the copied page. Or it will setup a readonly spte or readonly tlb mapping if it's a guest-read, if it calls get_user_pages with write=0. This is just an example. This allows to map any page pointed by any pte (and in turn visible in the primary CPU MMU), into a secondary MMU (be it a pure tlb like GRU, or an full MMU with both sptes and secondary-tlb like the shadow-pagetable layer with kvm), or a remote DMA in software like XPMEM (hence needing of schedule in XPMEM code to send the invalidate to the remote node, while no need to schedule in kvm/gru as it's an immediate event like invalidating primary-mmu pte). At least for KVM without this patch it's impossible to swap guests reliably. And having this feature and removing the page pin allows several other optimizations that simplify life considerably. Dependencies: 1) mm_take_all_locks() to register the mmu notifier when the whole VM isn't doing anything with "mm". This allows mmu notifier users to keep track if the VM is in the middle of the invalidate_range_begin/end critical section with an atomic counter incraese in range_begin and decreased in range_end. No secondary MMU page fault is allowed to map any spte or secondary tlb reference, while the VM is in the middle of range_begin/end as any page returned by get_user_pages in that critical section could later immediately be freed without any further ->invalidate_page notification (invalidate_range_begin/end works on ranges and ->invalidate_page isn't called immediately before freeing the page). To stop all page freeing and pagetable overwrites the mmap_sem must be taken in write mode and all other anon_vma/i_mmap locks must be taken too. 2) It'd be a waste to add branches in the VM if nobody could possibly run KVM/GRU/XPMEM on the kernel, so mmu notifiers will only enabled if CONFIG_KVM=m/y. In the current kernel kvm won't yet take advantage of mmu notifiers, but this already allows to compile a KVM external module against a kernel with mmu notifiers enabled and from the next pull from kvm.git we'll start using them. And GRU/XPMEM will also be able to continue the development by enabling KVM=m in their config, until they submit all GRU/XPMEM GPLv2 code to the mainline kernel. Then they can also enable MMU_NOTIFIERS in the same way KVM does it (even if KVM=n). This guarantees nobody selects MMU_NOTIFIER=y if KVM and GRU and XPMEM are all =n. The mmu_notifier_register call can fail because mm_take_all_locks may be interrupted by a signal and return -EINTR. Because mmu_notifier_reigster is used when a driver startup, a failure can be gracefully handled. Here an example of the change applied to kvm to register the mmu notifiers. Usually when a driver startups other allocations are required anyway and -ENOMEM failure paths exists already. struct kvm *kvm_arch_create_vm(void) { struct kvm *kvm = kzalloc(sizeof(struct kvm), GFP_KERNEL); + int err; if (!kvm) return ERR_PTR(-ENOMEM); INIT_LIST_HEAD(&kvm->arch.active_mmu_pages); + kvm->arch.mmu_notifier.ops = &kvm_mmu_notifier_ops; + err = mmu_notifier_register(&kvm->arch.mmu_notifier, current->mm); + if (err) { + kfree(kvm); + return ERR_PTR(err); + } + return kvm; } mmu_notifier_unregister returns void and it's reliable. The patch also adds a few needed but missing includes that would prevent kernel to compile after these changes on non-x86 archs (x86 didn't need them by luck). [akpm@linux-foundation.org: coding-style fixes] [akpm@linux-foundation.org: fix mm/filemap_xip.c build] [akpm@linux-foundation.org: fix mm/mmu_notifier.c build] Signed-off-by: Andrea Arcangeli <andrea@qumranet.com> Signed-off-by: Nick Piggin <npiggin@suse.de> Signed-off-by: Christoph Lameter <cl@linux-foundation.org> Cc: Jack Steiner <steiner@sgi.com> Cc: Robin Holt <holt@sgi.com> Cc: Nick Piggin <npiggin@suse.de> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: Kanoj Sarcar <kanojsarcar@yahoo.com> Cc: Roland Dreier <rdreier@cisco.com> Cc: Steve Wise <swise@opengridcomputing.com> Cc: Avi Kivity <avi@qumranet.com> Cc: Hugh Dickins <hugh@veritas.com> Cc: Rusty Russell <rusty@rustcorp.com.au> Cc: Anthony Liguori <aliguori@us.ibm.com> Cc: Chris Wright <chrisw@redhat.com> Cc: Marcelo Tosatti <marcelo@kvack.org> Cc: Eric Dumazet <dada1@cosmosbay.com> Cc: "Paul E. McKenney" <paulmck@us.ibm.com> Cc: Izik Eidus <izike@qumranet.com> Cc: Anthony Liguori <aliguori@us.ibm.com> Cc: Rik van Riel <riel@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-07-28 22:46:29 +00:00
mmu_notifier_mm_destroy(mm);
check_mm(mm);
mm: Add a user_ns owner to mm_struct and fix ptrace permission checks During exec dumpable is cleared if the file that is being executed is not readable by the user executing the file. A bug in ptrace_may_access allows reading the file if the executable happens to enter into a subordinate user namespace (aka clone(CLONE_NEWUSER), unshare(CLONE_NEWUSER), or setns(fd, CLONE_NEWUSER). This problem is fixed with only necessary userspace breakage by adding a user namespace owner to mm_struct, captured at the time of exec, so it is clear in which user namespace CAP_SYS_PTRACE must be present in to be able to safely give read permission to the executable. The function ptrace_may_access is modified to verify that the ptracer has CAP_SYS_ADMIN in task->mm->user_ns instead of task->cred->user_ns. This ensures that if the task changes it's cred into a subordinate user namespace it does not become ptraceable. The function ptrace_attach is modified to only set PT_PTRACE_CAP when CAP_SYS_PTRACE is held over task->mm->user_ns. The intent of PT_PTRACE_CAP is to be a flag to note that whatever permission changes the task might go through the tracer has sufficient permissions for it not to be an issue. task->cred->user_ns is always the same as or descendent of mm->user_ns. Which guarantees that having CAP_SYS_PTRACE over mm->user_ns is the worst case for the tasks credentials. To prevent regressions mm->dumpable and mm->user_ns are not considered when a task has no mm. As simply failing ptrace_may_attach causes regressions in privileged applications attempting to read things such as /proc/<pid>/stat Cc: stable@vger.kernel.org Acked-by: Kees Cook <keescook@chromium.org> Tested-by: Cyrill Gorcunov <gorcunov@openvz.org> Fixes: 8409cca70561 ("userns: allow ptrace from non-init user namespaces") Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2016-10-14 02:23:16 +00:00
put_user_ns(mm->user_ns);
free_mm(mm);
}
EXPORT_SYMBOL_GPL(__mmdrop);
static inline void __mmput(struct mm_struct *mm)
{
VM_BUG_ON(atomic_read(&mm->mm_users));
uprobe_clear_state(mm);
exit_aio(mm);
ksm_exit(mm);
khugepaged_exit(mm); /* must run before exit_mmap */
exit_mmap(mm);
thp: reduce usage of huge zero page's atomic counter The global zero page is used to satisfy an anonymous read fault. If THP(Transparent HugePage) is enabled then the global huge zero page is used. The global huge zero page uses an atomic counter for reference counting and is allocated/freed dynamically according to its counter value. CPU time spent on that counter will greatly increase if there are a lot of processes doing anonymous read faults. This patch proposes a way to reduce the access to the global counter so that the CPU load can be reduced accordingly. To do this, a new flag of the mm_struct is introduced: MMF_USED_HUGE_ZERO_PAGE. With this flag, the process only need to touch the global counter in two cases: 1 The first time it uses the global huge zero page; 2 The time when mm_user of its mm_struct reaches zero. Note that right now, the huge zero page is eligible to be freed as soon as its last use goes away. With this patch, the page will not be eligible to be freed until the exit of the last process from which it was ever used. And with the use of mm_user, the kthread is not eligible to use huge zero page either. Since no kthread is using huge zero page today, there is no difference after applying this patch. But if that is not desired, I can change it to when mm_count reaches zero. Case used for test on Haswell EP: usemem -n 72 --readonly -j 0x200000 100G Which spawns 72 processes and each will mmap 100G anonymous space and then do read only access to that space sequentially with a step of 2MB. CPU cycles from perf report for base commit: 54.03% usemem [kernel.kallsyms] [k] get_huge_zero_page CPU cycles from perf report for this commit: 0.11% usemem [kernel.kallsyms] [k] mm_get_huge_zero_page Performance(throughput) of the workload for base commit: 1784430792 Performance(throughput) of the workload for this commit: 4726928591 164% increase. Runtime of the workload for base commit: 707592 us Runtime of the workload for this commit: 303970 us 50% drop. Link: http://lkml.kernel.org/r/fe51a88f-446a-4622-1363-ad1282d71385@intel.com Signed-off-by: Aaron Lu <aaron.lu@intel.com> Cc: Sergey Senozhatsky <sergey.senozhatsky@gmail.com> Cc: "Kirill A. Shutemov" <kirill.shutemov@linux.intel.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: Tim Chen <tim.c.chen@linux.intel.com> Cc: Huang Ying <ying.huang@intel.com> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Jerome Marchand <jmarchan@redhat.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Mel Gorman <mgorman@techsingularity.net> Cc: Ebru Akagunduz <ebru.akagunduz@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-10-08 00:00:08 +00:00
mm_put_huge_zero_page(mm);
set_mm_exe_file(mm, NULL);
if (!list_empty(&mm->mmlist)) {
spin_lock(&mmlist_lock);
list_del(&mm->mmlist);
spin_unlock(&mmlist_lock);
}
if (mm->binfmt)
module_put(mm->binfmt->module);
set_bit(MMF_OOM_SKIP, &mm->flags);
mmdrop(mm);
}
/*
* Decrement the use count and release all resources for an mm.
*/
void mmput(struct mm_struct *mm)
{
might_sleep();
if (atomic_dec_and_test(&mm->mm_users))
__mmput(mm);
}
EXPORT_SYMBOL_GPL(mmput);
#ifdef CONFIG_MMU
static void mmput_async_fn(struct work_struct *work)
{
struct mm_struct *mm = container_of(work, struct mm_struct, async_put_work);
__mmput(mm);
}
void mmput_async(struct mm_struct *mm)
{
if (atomic_dec_and_test(&mm->mm_users)) {
INIT_WORK(&mm->async_put_work, mmput_async_fn);
schedule_work(&mm->async_put_work);
}
}
#endif
/**
* set_mm_exe_file - change a reference to the mm's executable file
*
* This changes mm's executable file (shown as symlink /proc/[pid]/exe).
*
* Main users are mmput() and sys_execve(). Callers prevent concurrent
* invocations: in mmput() nobody alive left, in execve task is single
* threaded. sys_prctl(PR_SET_MM_MAP/EXE_FILE) also needs to set the
* mm->exe_file, but does so without using set_mm_exe_file() in order
* to do avoid the need for any locks.
*/
void set_mm_exe_file(struct mm_struct *mm, struct file *new_exe_file)
{
struct file *old_exe_file;
/*
* It is safe to dereference the exe_file without RCU as
* this function is only called if nobody else can access
* this mm -- see comment above for justification.
*/
old_exe_file = rcu_dereference_raw(mm->exe_file);
if (new_exe_file)
get_file(new_exe_file);
rcu_assign_pointer(mm->exe_file, new_exe_file);
if (old_exe_file)
fput(old_exe_file);
}
/**
* get_mm_exe_file - acquire a reference to the mm's executable file
*
* Returns %NULL if mm has no associated executable file.
* User must release file via fput().
*/
struct file *get_mm_exe_file(struct mm_struct *mm)
{
struct file *exe_file;
rcu_read_lock();
exe_file = rcu_dereference(mm->exe_file);
if (exe_file && !get_file_rcu(exe_file))
exe_file = NULL;
rcu_read_unlock();
return exe_file;
}
EXPORT_SYMBOL(get_mm_exe_file);
/**
* get_task_exe_file - acquire a reference to the task's executable file
*
* Returns %NULL if task's mm (if any) has no associated executable file or
* this is a kernel thread with borrowed mm (see the comment above get_task_mm).
* User must release file via fput().
*/
struct file *get_task_exe_file(struct task_struct *task)
{
struct file *exe_file = NULL;
struct mm_struct *mm;
task_lock(task);
mm = task->mm;
if (mm) {
if (!(task->flags & PF_KTHREAD))
exe_file = get_mm_exe_file(mm);
}
task_unlock(task);
return exe_file;
}
EXPORT_SYMBOL(get_task_exe_file);
/**
* get_task_mm - acquire a reference to the task's mm
*
* Returns %NULL if the task has no mm. Checks PF_KTHREAD (meaning
* this kernel workthread has transiently adopted a user mm with use_mm,
* to do its AIO) is not set and if so returns a reference to it, after
* bumping up the use count. User must release the mm via mmput()
* after use. Typically used by /proc and ptrace.
*/
struct mm_struct *get_task_mm(struct task_struct *task)
{
struct mm_struct *mm;
task_lock(task);
mm = task->mm;
if (mm) {
if (task->flags & PF_KTHREAD)
mm = NULL;
else
mmget(mm);
}
task_unlock(task);
return mm;
}
EXPORT_SYMBOL_GPL(get_task_mm);
struct mm_struct *mm_access(struct task_struct *task, unsigned int mode)
{
struct mm_struct *mm;
int err;
err = mutex_lock_killable(&task->signal->cred_guard_mutex);
if (err)
return ERR_PTR(err);
mm = get_task_mm(task);
if (mm && mm != current->mm &&
!ptrace_may_access(task, mode)) {
mmput(mm);
mm = ERR_PTR(-EACCES);
}
mutex_unlock(&task->signal->cred_guard_mutex);
return mm;
}
static void complete_vfork_done(struct task_struct *tsk)
{
struct completion *vfork;
task_lock(tsk);
vfork = tsk->vfork_done;
if (likely(vfork)) {
tsk->vfork_done = NULL;
complete(vfork);
}
task_unlock(tsk);
}
static int wait_for_vfork_done(struct task_struct *child,
struct completion *vfork)
{
int killed;
freezer_do_not_count();
killed = wait_for_completion_killable(vfork);
freezer_count();
if (killed) {
task_lock(child);
child->vfork_done = NULL;
task_unlock(child);
}
put_task_struct(child);
return killed;
}
/* Please note the differences between mmput and mm_release.
* mmput is called whenever we stop holding onto a mm_struct,
* error success whatever.
*
* mm_release is called after a mm_struct has been removed
* from the current process.
*
* This difference is important for error handling, when we
* only half set up a mm_struct for a new process and need to restore
* the old one. Because we mmput the new mm_struct before
* restoring the old one. . .
* Eric Biederman 10 January 1998
*/
void mm_release(struct task_struct *tsk, struct mm_struct *mm)
{
/* Get rid of any futexes when releasing the mm */
#ifdef CONFIG_FUTEX
if (unlikely(tsk->robust_list)) {
exit_robust_list(tsk);
tsk->robust_list = NULL;
}
#ifdef CONFIG_COMPAT
if (unlikely(tsk->compat_robust_list)) {
compat_exit_robust_list(tsk);
tsk->compat_robust_list = NULL;
}
#endif
if (unlikely(!list_empty(&tsk->pi_state_list)))
exit_pi_state_list(tsk);
#endif
uprobes/core: Handle breakpoint and singlestep exceptions Uprobes uses exception notifiers to get to know if a thread hit a breakpoint or a singlestep exception. When a thread hits a uprobe or is singlestepping post a uprobe hit, the uprobe exception notifier sets its TIF_UPROBE bit, which will then be checked on its return to userspace path (do_notify_resume() ->uprobe_notify_resume()), where the consumers handlers are run (in task context) based on the defined filters. Uprobe hits are thread specific and hence we need to maintain information about if a task hit a uprobe, what uprobe was hit, the slot where the original instruction was copied for xol so that it can be singlestepped with appropriate fixups. In some cases, special care is needed for instructions that are executed out of line (xol). These are architecture specific artefacts, such as handling RIP relative instructions on x86_64. Since the instruction at which the uprobe was inserted is executed out of line, architecture specific fixups are added so that the thread continues normal execution in the presence of a uprobe. Postpone the signals until we execute the probed insn. post_xol() path does a recalc_sigpending() before return to user-mode, this ensures the signal can't be lost. Uprobes relies on DIE_DEBUG notification to notify if a singlestep is complete. Adds x86 specific uprobe exception notifiers and appropriate hooks needed to determine a uprobe hit and subsequent post processing. Add requisite x86 fixups for xol for uprobes. Specific cases needing fixups include relative jumps (x86_64), calls, etc. Where possible, we check and skip singlestepping the breakpointed instructions. For now we skip single byte as well as few multibyte nop instructions. However this can be extended to other instructions too. Credits to Oleg Nesterov for suggestions/patches related to signal, breakpoint, singlestep handling code. Signed-off-by: Srikar Dronamraju <srikar@linux.vnet.ibm.com> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Ananth N Mavinakayanahalli <ananth@in.ibm.com> Cc: Jim Keniston <jkenisto@linux.vnet.ibm.com> Cc: Linux-mm <linux-mm@kvack.org> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Andi Kleen <andi@firstfloor.org> Cc: Christoph Hellwig <hch@infradead.org> Cc: Steven Rostedt <rostedt@goodmis.org> Cc: Arnaldo Carvalho de Melo <acme@infradead.org> Cc: Masami Hiramatsu <masami.hiramatsu.pt@hitachi.com> Cc: Peter Zijlstra <peterz@infradead.org> Link: http://lkml.kernel.org/r/20120313180011.29771.89027.sendpatchset@srdronam.in.ibm.com [ Performed various cleanliness edits ] Signed-off-by: Ingo Molnar <mingo@elte.hu>
2012-03-13 18:00:11 +00:00
uprobe_free_utask(tsk);
/* Get rid of any cached register state */
deactivate_mm(tsk, mm);
/*
kernel/fork: fix CLONE_CHILD_CLEARTID regression in nscd Commit fec1d0115240 ("[PATCH] Disable CLONE_CHILD_CLEARTID for abnormal exit") has caused a subtle regression in nscd which uses CLONE_CHILD_CLEARTID to clear the nscd_certainly_running flag in the shared databases, so that the clients are notified when nscd is restarted. Now, when nscd uses a non-persistent database, clients that have it mapped keep thinking the database is being updated by nscd, when in fact nscd has created a new (anonymous) one (for non-persistent databases it uses an unlinked file as backend). The original proposal for the CLONE_CHILD_CLEARTID change claimed (https://lkml.org/lkml/2006/10/25/233): : The NPTL library uses the CLONE_CHILD_CLEARTID flag on clone() syscalls : on behalf of pthread_create() library calls. This feature is used to : request that the kernel clear the thread-id in user space (at an address : provided in the syscall) when the thread disassociates itself from the : address space, which is done in mm_release(). : : Unfortunately, when a multi-threaded process incurs a core dump (such as : from a SIGSEGV), the core-dumping thread sends SIGKILL signals to all of : the other threads, which then proceed to clear their user-space tids : before synchronizing in exit_mm() with the start of core dumping. This : misrepresents the state of process's address space at the time of the : SIGSEGV and makes it more difficult for someone to debug NPTL and glibc : problems (misleading him/her to conclude that the threads had gone away : before the fault). : : The fix below is to simply avoid the CLONE_CHILD_CLEARTID action if a : core dump has been initiated. The resulting patch from Roland (https://lkml.org/lkml/2006/10/26/269) seems to have a larger scope than the original patch asked for. It seems that limitting the scope of the check to core dumping should work for SIGSEGV issue describe above. [Changelog partly based on Andreas' description] Fixes: fec1d0115240 ("[PATCH] Disable CLONE_CHILD_CLEARTID for abnormal exit") Link: http://lkml.kernel.org/r/1471968749-26173-1-git-send-email-mhocko@kernel.org Signed-off-by: Michal Hocko <mhocko@suse.com> Tested-by: William Preston <wpreston@suse.com> Acked-by: Oleg Nesterov <oleg@redhat.com> Cc: Roland McGrath <roland@hack.frob.com> Cc: Andreas Schwab <schwab@suse.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-09-01 23:15:13 +00:00
* Signal userspace if we're not exiting with a core dump
* because we want to leave the value intact for debugging
* purposes.
*/
execve: must clear current->clear_child_tid While looking at Jens Rosenboom bug report (http://lkml.org/lkml/2009/7/27/35) about strange sys_futex call done from a dying "ps" program, we found following problem. clone() syscall has special support for TID of created threads. This support includes two features. One (CLONE_CHILD_SETTID) is to set an integer into user memory with the TID value. One (CLONE_CHILD_CLEARTID) is to clear this same integer once the created thread dies. The integer location is a user provided pointer, provided at clone() time. kernel keeps this pointer value into current->clear_child_tid. At execve() time, we should make sure kernel doesnt keep this user provided pointer, as full user memory is replaced by a new one. As glibc fork() actually uses clone() syscall with CLONE_CHILD_SETTID and CLONE_CHILD_CLEARTID set, chances are high that we might corrupt user memory in forked processes. Following sequence could happen: 1) bash (or any program) starts a new process, by a fork() call that glibc maps to a clone( ... CLONE_CHILD_SETTID | CLONE_CHILD_CLEARTID ...) syscall 2) When new process starts, its current->clear_child_tid is set to a location that has a meaning only in bash (or initial program) context (&THREAD_SELF->tid) 3) This new process does the execve() syscall to start a new program. current->clear_child_tid is left unchanged (a non NULL value) 4) If this new program creates some threads, and initial thread exits, kernel will attempt to clear the integer pointed by current->clear_child_tid from mm_release() : if (tsk->clear_child_tid && !(tsk->flags & PF_SIGNALED) && atomic_read(&mm->mm_users) > 1) { u32 __user * tidptr = tsk->clear_child_tid; tsk->clear_child_tid = NULL; /* * We don't check the error code - if userspace has * not set up a proper pointer then tough luck. */ << here >> put_user(0, tidptr); sys_futex(tidptr, FUTEX_WAKE, 1, NULL, NULL, 0); } 5) OR : if new program is not multi-threaded, but spied by /proc/pid users (ps command for example), mm_users > 1, and the exiting program could corrupt 4 bytes in a persistent memory area (shm or memory mapped file) If current->clear_child_tid points to a writeable portion of memory of the new program, kernel happily and silently corrupts 4 bytes of memory, with unexpected effects. Fix is straightforward and should not break any sane program. Reported-by: Jens Rosenboom <jens@mcbone.net> Acked-by: Linus Torvalds <torvalds@linux-foundation.org> Signed-off-by: Eric Dumazet <eric.dumazet@gmail.com> Signed-off-by: Oleg Nesterov <oleg@redhat.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Sonny Rao <sonnyrao@us.ibm.com> Cc: Ingo Molnar <mingo@elte.hu> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ulrich Drepper <drepper@redhat.com> Cc: Oleg Nesterov <oleg@redhat.com> Cc: <stable@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-08-06 22:09:28 +00:00
if (tsk->clear_child_tid) {
kernel/fork: fix CLONE_CHILD_CLEARTID regression in nscd Commit fec1d0115240 ("[PATCH] Disable CLONE_CHILD_CLEARTID for abnormal exit") has caused a subtle regression in nscd which uses CLONE_CHILD_CLEARTID to clear the nscd_certainly_running flag in the shared databases, so that the clients are notified when nscd is restarted. Now, when nscd uses a non-persistent database, clients that have it mapped keep thinking the database is being updated by nscd, when in fact nscd has created a new (anonymous) one (for non-persistent databases it uses an unlinked file as backend). The original proposal for the CLONE_CHILD_CLEARTID change claimed (https://lkml.org/lkml/2006/10/25/233): : The NPTL library uses the CLONE_CHILD_CLEARTID flag on clone() syscalls : on behalf of pthread_create() library calls. This feature is used to : request that the kernel clear the thread-id in user space (at an address : provided in the syscall) when the thread disassociates itself from the : address space, which is done in mm_release(). : : Unfortunately, when a multi-threaded process incurs a core dump (such as : from a SIGSEGV), the core-dumping thread sends SIGKILL signals to all of : the other threads, which then proceed to clear their user-space tids : before synchronizing in exit_mm() with the start of core dumping. This : misrepresents the state of process's address space at the time of the : SIGSEGV and makes it more difficult for someone to debug NPTL and glibc : problems (misleading him/her to conclude that the threads had gone away : before the fault). : : The fix below is to simply avoid the CLONE_CHILD_CLEARTID action if a : core dump has been initiated. The resulting patch from Roland (https://lkml.org/lkml/2006/10/26/269) seems to have a larger scope than the original patch asked for. It seems that limitting the scope of the check to core dumping should work for SIGSEGV issue describe above. [Changelog partly based on Andreas' description] Fixes: fec1d0115240 ("[PATCH] Disable CLONE_CHILD_CLEARTID for abnormal exit") Link: http://lkml.kernel.org/r/1471968749-26173-1-git-send-email-mhocko@kernel.org Signed-off-by: Michal Hocko <mhocko@suse.com> Tested-by: William Preston <wpreston@suse.com> Acked-by: Oleg Nesterov <oleg@redhat.com> Cc: Roland McGrath <roland@hack.frob.com> Cc: Andreas Schwab <schwab@suse.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-09-01 23:15:13 +00:00
if (!(tsk->signal->flags & SIGNAL_GROUP_COREDUMP) &&
execve: must clear current->clear_child_tid While looking at Jens Rosenboom bug report (http://lkml.org/lkml/2009/7/27/35) about strange sys_futex call done from a dying "ps" program, we found following problem. clone() syscall has special support for TID of created threads. This support includes two features. One (CLONE_CHILD_SETTID) is to set an integer into user memory with the TID value. One (CLONE_CHILD_CLEARTID) is to clear this same integer once the created thread dies. The integer location is a user provided pointer, provided at clone() time. kernel keeps this pointer value into current->clear_child_tid. At execve() time, we should make sure kernel doesnt keep this user provided pointer, as full user memory is replaced by a new one. As glibc fork() actually uses clone() syscall with CLONE_CHILD_SETTID and CLONE_CHILD_CLEARTID set, chances are high that we might corrupt user memory in forked processes. Following sequence could happen: 1) bash (or any program) starts a new process, by a fork() call that glibc maps to a clone( ... CLONE_CHILD_SETTID | CLONE_CHILD_CLEARTID ...) syscall 2) When new process starts, its current->clear_child_tid is set to a location that has a meaning only in bash (or initial program) context (&THREAD_SELF->tid) 3) This new process does the execve() syscall to start a new program. current->clear_child_tid is left unchanged (a non NULL value) 4) If this new program creates some threads, and initial thread exits, kernel will attempt to clear the integer pointed by current->clear_child_tid from mm_release() : if (tsk->clear_child_tid && !(tsk->flags & PF_SIGNALED) && atomic_read(&mm->mm_users) > 1) { u32 __user * tidptr = tsk->clear_child_tid; tsk->clear_child_tid = NULL; /* * We don't check the error code - if userspace has * not set up a proper pointer then tough luck. */ << here >> put_user(0, tidptr); sys_futex(tidptr, FUTEX_WAKE, 1, NULL, NULL, 0); } 5) OR : if new program is not multi-threaded, but spied by /proc/pid users (ps command for example), mm_users > 1, and the exiting program could corrupt 4 bytes in a persistent memory area (shm or memory mapped file) If current->clear_child_tid points to a writeable portion of memory of the new program, kernel happily and silently corrupts 4 bytes of memory, with unexpected effects. Fix is straightforward and should not break any sane program. Reported-by: Jens Rosenboom <jens@mcbone.net> Acked-by: Linus Torvalds <torvalds@linux-foundation.org> Signed-off-by: Eric Dumazet <eric.dumazet@gmail.com> Signed-off-by: Oleg Nesterov <oleg@redhat.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Sonny Rao <sonnyrao@us.ibm.com> Cc: Ingo Molnar <mingo@elte.hu> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ulrich Drepper <drepper@redhat.com> Cc: Oleg Nesterov <oleg@redhat.com> Cc: <stable@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-08-06 22:09:28 +00:00
atomic_read(&mm->mm_users) > 1) {
/*
* We don't check the error code - if userspace has
* not set up a proper pointer then tough luck.
*/
put_user(0, tsk->clear_child_tid);
sys_futex(tsk->clear_child_tid, FUTEX_WAKE,
1, NULL, NULL, 0);
}
tsk->clear_child_tid = NULL;
}
/*
* All done, finally we can wake up parent and return this mm to him.
* Also kthread_stop() uses this completion for synchronization.
*/
if (tsk->vfork_done)
complete_vfork_done(tsk);
}
/*
* Allocate a new mm structure and copy contents from the
* mm structure of the passed in task structure.
*/
static struct mm_struct *dup_mm(struct task_struct *tsk)
{
struct mm_struct *mm, *oldmm = current->mm;
int err;
mm = allocate_mm();
if (!mm)
goto fail_nomem;
memcpy(mm, oldmm, sizeof(*mm));
mm: Add a user_ns owner to mm_struct and fix ptrace permission checks During exec dumpable is cleared if the file that is being executed is not readable by the user executing the file. A bug in ptrace_may_access allows reading the file if the executable happens to enter into a subordinate user namespace (aka clone(CLONE_NEWUSER), unshare(CLONE_NEWUSER), or setns(fd, CLONE_NEWUSER). This problem is fixed with only necessary userspace breakage by adding a user namespace owner to mm_struct, captured at the time of exec, so it is clear in which user namespace CAP_SYS_PTRACE must be present in to be able to safely give read permission to the executable. The function ptrace_may_access is modified to verify that the ptracer has CAP_SYS_ADMIN in task->mm->user_ns instead of task->cred->user_ns. This ensures that if the task changes it's cred into a subordinate user namespace it does not become ptraceable. The function ptrace_attach is modified to only set PT_PTRACE_CAP when CAP_SYS_PTRACE is held over task->mm->user_ns. The intent of PT_PTRACE_CAP is to be a flag to note that whatever permission changes the task might go through the tracer has sufficient permissions for it not to be an issue. task->cred->user_ns is always the same as or descendent of mm->user_ns. Which guarantees that having CAP_SYS_PTRACE over mm->user_ns is the worst case for the tasks credentials. To prevent regressions mm->dumpable and mm->user_ns are not considered when a task has no mm. As simply failing ptrace_may_attach causes regressions in privileged applications attempting to read things such as /proc/<pid>/stat Cc: stable@vger.kernel.org Acked-by: Kees Cook <keescook@chromium.org> Tested-by: Cyrill Gorcunov <gorcunov@openvz.org> Fixes: 8409cca70561 ("userns: allow ptrace from non-init user namespaces") Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2016-10-14 02:23:16 +00:00
if (!mm_init(mm, tsk, mm->user_ns))
goto fail_nomem;
err = dup_mmap(mm, oldmm);
if (err)
goto free_pt;
mm->hiwater_rss = get_mm_rss(mm);
mm->hiwater_vm = mm->total_vm;
if (mm->binfmt && !try_module_get(mm->binfmt->module))
goto free_pt;
return mm;
free_pt:
/* don't put binfmt in mmput, we haven't got module yet */
mm->binfmt = NULL;
mmput(mm);
fail_nomem:
return NULL;
}
static int copy_mm(unsigned long clone_flags, struct task_struct *tsk)
{
struct mm_struct *mm, *oldmm;
int retval;
tsk->min_flt = tsk->maj_flt = 0;
tsk->nvcsw = tsk->nivcsw = 0;
#ifdef CONFIG_DETECT_HUNG_TASK
tsk->last_switch_count = tsk->nvcsw + tsk->nivcsw;
#endif
tsk->mm = NULL;
tsk->active_mm = NULL;
/*
* Are we cloning a kernel thread?
*
* We need to steal a active VM for that..
*/
oldmm = current->mm;
if (!oldmm)
return 0;
mm: per-thread vma caching This patch is a continuation of efforts trying to optimize find_vma(), avoiding potentially expensive rbtree walks to locate a vma upon faults. The original approach (https://lkml.org/lkml/2013/11/1/410), where the largest vma was also cached, ended up being too specific and random, thus further comparison with other approaches were needed. There are two things to consider when dealing with this, the cache hit rate and the latency of find_vma(). Improving the hit-rate does not necessarily translate in finding the vma any faster, as the overhead of any fancy caching schemes can be too high to consider. We currently cache the last used vma for the whole address space, which provides a nice optimization, reducing the total cycles in find_vma() by up to 250%, for workloads with good locality. On the other hand, this simple scheme is pretty much useless for workloads with poor locality. Analyzing ebizzy runs shows that, no matter how many threads are running, the mmap_cache hit rate is less than 2%, and in many situations below 1%. The proposed approach is to replace this scheme with a small per-thread cache, maximizing hit rates at a very low maintenance cost. Invalidations are performed by simply bumping up a 32-bit sequence number. The only expensive operation is in the rare case of a seq number overflow, where all caches that share the same address space are flushed. Upon a miss, the proposed replacement policy is based on the page number that contains the virtual address in question. Concretely, the following results are seen on an 80 core, 8 socket x86-64 box: 1) System bootup: Most programs are single threaded, so the per-thread scheme does improve ~50% hit rate by just adding a few more slots to the cache. +----------------+----------+------------------+ | caching scheme | hit-rate | cycles (billion) | +----------------+----------+------------------+ | baseline | 50.61% | 19.90 | | patched | 73.45% | 13.58 | +----------------+----------+------------------+ 2) Kernel build: This one is already pretty good with the current approach as we're dealing with good locality. +----------------+----------+------------------+ | caching scheme | hit-rate | cycles (billion) | +----------------+----------+------------------+ | baseline | 75.28% | 11.03 | | patched | 88.09% | 9.31 | +----------------+----------+------------------+ 3) Oracle 11g Data Mining (4k pages): Similar to the kernel build workload. +----------------+----------+------------------+ | caching scheme | hit-rate | cycles (billion) | +----------------+----------+------------------+ | baseline | 70.66% | 17.14 | | patched | 91.15% | 12.57 | +----------------+----------+------------------+ 4) Ebizzy: There's a fair amount of variation from run to run, but this approach always shows nearly perfect hit rates, while baseline is just about non-existent. The amounts of cycles can fluctuate between anywhere from ~60 to ~116 for the baseline scheme, but this approach reduces it considerably. For instance, with 80 threads: +----------------+----------+------------------+ | caching scheme | hit-rate | cycles (billion) | +----------------+----------+------------------+ | baseline | 1.06% | 91.54 | | patched | 99.97% | 14.18 | +----------------+----------+------------------+ [akpm@linux-foundation.org: fix nommu build, per Davidlohr] [akpm@linux-foundation.org: document vmacache_valid() logic] [akpm@linux-foundation.org: attempt to untangle header files] [akpm@linux-foundation.org: add vmacache_find() BUG_ON] [hughd@google.com: add vmacache_valid_mm() (from Oleg)] [akpm@linux-foundation.org: coding-style fixes] [akpm@linux-foundation.org: adjust and enhance comments] Signed-off-by: Davidlohr Bueso <davidlohr@hp.com> Reviewed-by: Rik van Riel <riel@redhat.com> Acked-by: Linus Torvalds <torvalds@linux-foundation.org> Reviewed-by: Michel Lespinasse <walken@google.com> Cc: Oleg Nesterov <oleg@redhat.com> Tested-by: Hugh Dickins <hughd@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-04-07 22:37:25 +00:00
/* initialize the new vmacache entries */
vmacache_flush(tsk);
if (clone_flags & CLONE_VM) {
mmget(oldmm);
mm = oldmm;
goto good_mm;
}
retval = -ENOMEM;
mm = dup_mm(tsk);
if (!mm)
goto fail_nomem;
good_mm:
tsk->mm = mm;
tsk->active_mm = mm;
return 0;
fail_nomem:
return retval;
}
static int copy_fs(unsigned long clone_flags, struct task_struct *tsk)
{
struct fs_struct *fs = current->fs;
if (clone_flags & CLONE_FS) {
/* tsk->fs is already what we want */
spin_lock(&fs->lock);
if (fs->in_exec) {
spin_unlock(&fs->lock);
return -EAGAIN;
}
fs->users++;
spin_unlock(&fs->lock);
return 0;
}
tsk->fs = copy_fs_struct(fs);
if (!tsk->fs)
return -ENOMEM;
return 0;
}
static int copy_files(unsigned long clone_flags, struct task_struct *tsk)
{
struct files_struct *oldf, *newf;
int error = 0;
/*
* A background process may not have any files ...
*/
oldf = current->files;
if (!oldf)
goto out;
if (clone_flags & CLONE_FILES) {
atomic_inc(&oldf->count);
goto out;
}
newf = dup_fd(oldf, &error);
if (!newf)
goto out;
tsk->files = newf;
error = 0;
out:
return error;
}
static int copy_io(unsigned long clone_flags, struct task_struct *tsk)
{
#ifdef CONFIG_BLOCK
struct io_context *ioc = current->io_context;
block: make ioc get/put interface more conventional and fix race on alloction Ignoring copy_io() during fork, io_context can be allocated from two places - current_io_context() and set_task_ioprio(). The former is always called from local task while the latter can be called from different task. The synchornization between them are peculiar and dubious. * current_io_context() doesn't grab task_lock() and assumes that if it saw %NULL ->io_context, it would stay that way until allocation and assignment is complete. It has smp_wmb() between alloc/init and assignment. * set_task_ioprio() grabs task_lock() for assignment and does smp_read_barrier_depends() between "ioc = task->io_context" and "if (ioc)". Unfortunately, this doesn't achieve anything - the latter is not a dependent load of the former. ie, if ioc itself were being dereferenced "ioc->xxx", it would mean something (not sure what tho) but as the code currently stands, the dependent read barrier is noop. As only one of the the two test-assignment sequences is task_lock() protected, the task_lock() can't do much about race between the two. Nothing prevents current_io_context() and set_task_ioprio() allocating its own ioc for the same task and overwriting the other's. Also, set_task_ioprio() can race with exiting task and create a new ioc after exit_io_context() is finished. ioc get/put doesn't have any reason to be complex. The only hot path is accessing the existing ioc of %current, which is simple to achieve given that ->io_context is never destroyed as long as the task is alive. All other paths can happily go through task_lock() like all other task sub structures without impacting anything. This patch updates ioc get/put so that it becomes more conventional. * alloc_io_context() is replaced with get_task_io_context(). This is the only interface which can acquire access to ioc of another task. On return, the caller has an explicit reference to the object which should be put using put_io_context() afterwards. * The functionality of current_io_context() remains the same but when creating a new ioc, it shares the code path with get_task_io_context() and always goes through task_lock(). * get_io_context() now means incrementing ref on an ioc which the caller already has access to (be that an explicit refcnt or implicit %current one). * PF_EXITING inhibits creation of new io_context and once exit_io_context() is finished, it's guaranteed that both ioc acquisition functions return %NULL. * All users are updated. Most are trivial but smp_read_barrier_depends() removal from cfq_get_io_context() needs a bit of explanation. I suppose the original intention was to ensure ioc->ioprio is visible when set_task_ioprio() allocates new io_context and installs it; however, this wouldn't have worked because set_task_ioprio() doesn't have wmb between init and install. There are other problems with this which will be fixed in another patch. * While at it, use NUMA_NO_NODE instead of -1 for wildcard node specification. -v2: Vivek spotted contamination from debug patch. Removed. Signed-off-by: Tejun Heo <tj@kernel.org> Cc: Vivek Goyal <vgoyal@redhat.com> Signed-off-by: Jens Axboe <axboe@kernel.dk>
2011-12-13 23:33:38 +00:00
struct io_context *new_ioc;
if (!ioc)
return 0;
/*
* Share io context with parent, if CLONE_IO is set
*/
if (clone_flags & CLONE_IO) {
ioc_task_link(ioc);
tsk->io_context = ioc;
} else if (ioprio_valid(ioc->ioprio)) {
block: make ioc get/put interface more conventional and fix race on alloction Ignoring copy_io() during fork, io_context can be allocated from two places - current_io_context() and set_task_ioprio(). The former is always called from local task while the latter can be called from different task. The synchornization between them are peculiar and dubious. * current_io_context() doesn't grab task_lock() and assumes that if it saw %NULL ->io_context, it would stay that way until allocation and assignment is complete. It has smp_wmb() between alloc/init and assignment. * set_task_ioprio() grabs task_lock() for assignment and does smp_read_barrier_depends() between "ioc = task->io_context" and "if (ioc)". Unfortunately, this doesn't achieve anything - the latter is not a dependent load of the former. ie, if ioc itself were being dereferenced "ioc->xxx", it would mean something (not sure what tho) but as the code currently stands, the dependent read barrier is noop. As only one of the the two test-assignment sequences is task_lock() protected, the task_lock() can't do much about race between the two. Nothing prevents current_io_context() and set_task_ioprio() allocating its own ioc for the same task and overwriting the other's. Also, set_task_ioprio() can race with exiting task and create a new ioc after exit_io_context() is finished. ioc get/put doesn't have any reason to be complex. The only hot path is accessing the existing ioc of %current, which is simple to achieve given that ->io_context is never destroyed as long as the task is alive. All other paths can happily go through task_lock() like all other task sub structures without impacting anything. This patch updates ioc get/put so that it becomes more conventional. * alloc_io_context() is replaced with get_task_io_context(). This is the only interface which can acquire access to ioc of another task. On return, the caller has an explicit reference to the object which should be put using put_io_context() afterwards. * The functionality of current_io_context() remains the same but when creating a new ioc, it shares the code path with get_task_io_context() and always goes through task_lock(). * get_io_context() now means incrementing ref on an ioc which the caller already has access to (be that an explicit refcnt or implicit %current one). * PF_EXITING inhibits creation of new io_context and once exit_io_context() is finished, it's guaranteed that both ioc acquisition functions return %NULL. * All users are updated. Most are trivial but smp_read_barrier_depends() removal from cfq_get_io_context() needs a bit of explanation. I suppose the original intention was to ensure ioc->ioprio is visible when set_task_ioprio() allocates new io_context and installs it; however, this wouldn't have worked because set_task_ioprio() doesn't have wmb between init and install. There are other problems with this which will be fixed in another patch. * While at it, use NUMA_NO_NODE instead of -1 for wildcard node specification. -v2: Vivek spotted contamination from debug patch. Removed. Signed-off-by: Tejun Heo <tj@kernel.org> Cc: Vivek Goyal <vgoyal@redhat.com> Signed-off-by: Jens Axboe <axboe@kernel.dk>
2011-12-13 23:33:38 +00:00
new_ioc = get_task_io_context(tsk, GFP_KERNEL, NUMA_NO_NODE);
if (unlikely(!new_ioc))
return -ENOMEM;
block: make ioc get/put interface more conventional and fix race on alloction Ignoring copy_io() during fork, io_context can be allocated from two places - current_io_context() and set_task_ioprio(). The former is always called from local task while the latter can be called from different task. The synchornization between them are peculiar and dubious. * current_io_context() doesn't grab task_lock() and assumes that if it saw %NULL ->io_context, it would stay that way until allocation and assignment is complete. It has smp_wmb() between alloc/init and assignment. * set_task_ioprio() grabs task_lock() for assignment and does smp_read_barrier_depends() between "ioc = task->io_context" and "if (ioc)". Unfortunately, this doesn't achieve anything - the latter is not a dependent load of the former. ie, if ioc itself were being dereferenced "ioc->xxx", it would mean something (not sure what tho) but as the code currently stands, the dependent read barrier is noop. As only one of the the two test-assignment sequences is task_lock() protected, the task_lock() can't do much about race between the two. Nothing prevents current_io_context() and set_task_ioprio() allocating its own ioc for the same task and overwriting the other's. Also, set_task_ioprio() can race with exiting task and create a new ioc after exit_io_context() is finished. ioc get/put doesn't have any reason to be complex. The only hot path is accessing the existing ioc of %current, which is simple to achieve given that ->io_context is never destroyed as long as the task is alive. All other paths can happily go through task_lock() like all other task sub structures without impacting anything. This patch updates ioc get/put so that it becomes more conventional. * alloc_io_context() is replaced with get_task_io_context(). This is the only interface which can acquire access to ioc of another task. On return, the caller has an explicit reference to the object which should be put using put_io_context() afterwards. * The functionality of current_io_context() remains the same but when creating a new ioc, it shares the code path with get_task_io_context() and always goes through task_lock(). * get_io_context() now means incrementing ref on an ioc which the caller already has access to (be that an explicit refcnt or implicit %current one). * PF_EXITING inhibits creation of new io_context and once exit_io_context() is finished, it's guaranteed that both ioc acquisition functions return %NULL. * All users are updated. Most are trivial but smp_read_barrier_depends() removal from cfq_get_io_context() needs a bit of explanation. I suppose the original intention was to ensure ioc->ioprio is visible when set_task_ioprio() allocates new io_context and installs it; however, this wouldn't have worked because set_task_ioprio() doesn't have wmb between init and install. There are other problems with this which will be fixed in another patch. * While at it, use NUMA_NO_NODE instead of -1 for wildcard node specification. -v2: Vivek spotted contamination from debug patch. Removed. Signed-off-by: Tejun Heo <tj@kernel.org> Cc: Vivek Goyal <vgoyal@redhat.com> Signed-off-by: Jens Axboe <axboe@kernel.dk>
2011-12-13 23:33:38 +00:00
new_ioc->ioprio = ioc->ioprio;
put_io_context(new_ioc);
}
#endif
return 0;
}
static int copy_sighand(unsigned long clone_flags, struct task_struct *tsk)
{
struct sighand_struct *sig;
if (clone_flags & CLONE_SIGHAND) {
atomic_inc(&current->sighand->count);
return 0;
}
sig = kmem_cache_alloc(sighand_cachep, GFP_KERNEL);
rcu_assign_pointer(tsk->sighand, sig);
if (!sig)
return -ENOMEM;
atomic_set(&sig->count, 1);
memcpy(sig->action, current->sighand->action, sizeof(sig->action));
return 0;
}
void __cleanup_sighand(struct sighand_struct *sighand)
{
epoll: introduce POLLFREE to flush ->signalfd_wqh before kfree() This patch is intentionally incomplete to simplify the review. It ignores ep_unregister_pollwait() which plays with the same wqh. See the next change. epoll assumes that the EPOLL_CTL_ADD'ed file controls everything f_op->poll() needs. In particular it assumes that the wait queue can't go away until eventpoll_release(). This is not true in case of signalfd, the task which does EPOLL_CTL_ADD uses its ->sighand which is not connected to the file. This patch adds the special event, POLLFREE, currently only for epoll. It expects that init_poll_funcptr()'ed hook should do the necessary cleanup. Perhaps it should be defined as EPOLLFREE in eventpoll. __cleanup_sighand() is changed to do wake_up_poll(POLLFREE) if ->signalfd_wqh is not empty, we add the new signalfd_cleanup() helper. ep_poll_callback(POLLFREE) simply does list_del_init(task_list). This make this poll entry inconsistent, but we don't care. If you share epoll fd which contains our sigfd with another process you should blame yourself. signalfd is "really special". I simply do not know how we can define the "right" semantics if it used with epoll. The main problem is, epoll calls signalfd_poll() once to establish the connection with the wait queue, after that signalfd_poll(NULL) returns the different/inconsistent results depending on who does EPOLL_CTL_MOD/signalfd_read/etc. IOW: apart from sigmask, signalfd has nothing to do with the file, it works with the current thread. In short: this patch is the hack which tries to fix the symptoms. It also assumes that nobody can take tasklist_lock under epoll locks, this seems to be true. Note: - we do not have wake_up_all_poll() but wake_up_poll() is fine, poll/epoll doesn't use WQ_FLAG_EXCLUSIVE. - signalfd_cleanup() uses POLLHUP along with POLLFREE, we need a couple of simple changes in eventpoll.c to make sure it can't be "lost". Reported-by: Maxime Bizon <mbizon@freebox.fr> Cc: <stable@kernel.org> Signed-off-by: Oleg Nesterov <oleg@redhat.com> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-02-24 19:07:11 +00:00
if (atomic_dec_and_test(&sighand->count)) {
signalfd_cleanup(sighand);
/*
* sighand_cachep is SLAB_DESTROY_BY_RCU so we can free it
* without an RCU grace period, see __lock_task_sighand().
*/
kmem_cache_free(sighand_cachep, sighand);
epoll: introduce POLLFREE to flush ->signalfd_wqh before kfree() This patch is intentionally incomplete to simplify the review. It ignores ep_unregister_pollwait() which plays with the same wqh. See the next change. epoll assumes that the EPOLL_CTL_ADD'ed file controls everything f_op->poll() needs. In particular it assumes that the wait queue can't go away until eventpoll_release(). This is not true in case of signalfd, the task which does EPOLL_CTL_ADD uses its ->sighand which is not connected to the file. This patch adds the special event, POLLFREE, currently only for epoll. It expects that init_poll_funcptr()'ed hook should do the necessary cleanup. Perhaps it should be defined as EPOLLFREE in eventpoll. __cleanup_sighand() is changed to do wake_up_poll(POLLFREE) if ->signalfd_wqh is not empty, we add the new signalfd_cleanup() helper. ep_poll_callback(POLLFREE) simply does list_del_init(task_list). This make this poll entry inconsistent, but we don't care. If you share epoll fd which contains our sigfd with another process you should blame yourself. signalfd is "really special". I simply do not know how we can define the "right" semantics if it used with epoll. The main problem is, epoll calls signalfd_poll() once to establish the connection with the wait queue, after that signalfd_poll(NULL) returns the different/inconsistent results depending on who does EPOLL_CTL_MOD/signalfd_read/etc. IOW: apart from sigmask, signalfd has nothing to do with the file, it works with the current thread. In short: this patch is the hack which tries to fix the symptoms. It also assumes that nobody can take tasklist_lock under epoll locks, this seems to be true. Note: - we do not have wake_up_all_poll() but wake_up_poll() is fine, poll/epoll doesn't use WQ_FLAG_EXCLUSIVE. - signalfd_cleanup() uses POLLHUP along with POLLFREE, we need a couple of simple changes in eventpoll.c to make sure it can't be "lost". Reported-by: Maxime Bizon <mbizon@freebox.fr> Cc: <stable@kernel.org> Signed-off-by: Oleg Nesterov <oleg@redhat.com> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-02-24 19:07:11 +00:00
}
}
#ifdef CONFIG_POSIX_TIMERS
timers: fix itimer/many thread hang Overview This patch reworks the handling of POSIX CPU timers, including the ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together with the help of Roland McGrath, the owner and original writer of this code. The problem we ran into, and the reason for this rework, has to do with using a profiling timer in a process with a large number of threads. It appears that the performance of the old implementation of run_posix_cpu_timers() was at least O(n*3) (where "n" is the number of threads in a process) or worse. Everything is fine with an increasing number of threads until the time taken for that routine to run becomes the same as or greater than the tick time, at which point things degrade rather quickly. This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF." Code Changes This rework corrects the implementation of run_posix_cpu_timers() to make it run in constant time for a particular machine. (Performance may vary between one machine and another depending upon whether the kernel is built as single- or multiprocessor and, in the latter case, depending upon the number of running processors.) To do this, at each tick we now update fields in signal_struct as well as task_struct. The run_posix_cpu_timers() function uses those fields to make its decisions. We define a new structure, "task_cputime," to contain user, system and scheduler times and use these in appropriate places: struct task_cputime { cputime_t utime; cputime_t stime; unsigned long long sum_exec_runtime; }; This is included in the structure "thread_group_cputime," which is a new substructure of signal_struct and which varies for uniprocessor versus multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as a simple substructure, while for multiprocessor kernels it is a pointer: struct thread_group_cputime { struct task_cputime totals; }; struct thread_group_cputime { struct task_cputime *totals; }; We also add a new task_cputime substructure directly to signal_struct, to cache the earliest expiration of process-wide timers, and task_cputime also replaces the it_*_expires fields of task_struct (used for earliest expiration of thread timers). The "thread_group_cputime" structure contains process-wide timers that are updated via account_user_time() and friends. In the non-SMP case the structure is a simple aggregator; unfortunately in the SMP case that simplicity was not achievable due to cache-line contention between CPUs (in one measured case performance was actually _worse_ on a 16-cpu system than the same test on a 4-cpu system, due to this contention). For SMP, the thread_group_cputime counters are maintained as a per-cpu structure allocated using alloc_percpu(). The timer functions update only the timer field in the structure corresponding to the running CPU, obtained using per_cpu_ptr(). We define a set of inline functions in sched.h that we use to maintain the thread_group_cputime structure and hide the differences between UP and SMP implementations from the rest of the kernel. The thread_group_cputime_init() function initializes the thread_group_cputime structure for the given task. The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the out-of-line function thread_group_cputime_alloc_smp() to allocate and fill in the per-cpu structures and fields. The thread_group_cputime_free() function, also a no-op for UP, in SMP frees the per-cpu structures. The thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls thread_group_cputime_alloc() if the per-cpu structures haven't yet been allocated. The thread_group_cputime() function fills the task_cputime structure it is passed with the contents of the thread_group_cputime fields; in UP it's that simple but in SMP it must also safely check that tsk->signal is non-NULL (if it is it just uses the appropriate fields of task_struct) and, if so, sums the per-cpu values for each online CPU. Finally, the three functions account_group_user_time(), account_group_system_time() and account_group_exec_runtime() are used by timer functions to update the respective fields of the thread_group_cputime structure. Non-SMP operation is trivial and will not be mentioned further. The per-cpu structure is always allocated when a task creates its first new thread, via a call to thread_group_cputime_clone_thread() from copy_signal(). It is freed at process exit via a call to thread_group_cputime_free() from cleanup_signal(). All functions that formerly summed utime/stime/sum_sched_runtime values from from all threads in the thread group now use thread_group_cputime() to snapshot the values in the thread_group_cputime structure or the values in the task structure itself if the per-cpu structure hasn't been allocated. Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit. The run_posix_cpu_timers() function has been split into a fast path and a slow path; the former safely checks whether there are any expired thread timers and, if not, just returns, while the slow path does the heavy lifting. With the dedicated thread group fields, timers are no longer "rebalanced" and the process_timer_rebalance() function and related code has gone away. All summing loops are gone and all code that used them now uses the thread_group_cputime() inline. When process-wide timers are set, the new task_cputime structure in signal_struct is used to cache the earliest expiration; this is checked in the fast path. Performance The fix appears not to add significant overhead to existing operations. It generally performs the same as the current code except in two cases, one in which it performs slightly worse (Case 5 below) and one in which it performs very significantly better (Case 2 below). Overall it's a wash except in those two cases. I've since done somewhat more involved testing on a dual-core Opteron system. Case 1: With no itimer running, for a test with 100,000 threads, the fixed kernel took 1428.5 seconds, 513 seconds more than the unfixed system, all of which was spent in the system. There were twice as many voluntary context switches with the fix as without it. Case 2: With an itimer running at .01 second ticks and 4000 threads (the most an unmodified kernel can handle), the fixed kernel ran the test in eight percent of the time (5.8 seconds as opposed to 70 seconds) and had better tick accuracy (.012 seconds per tick as opposed to .023 seconds per tick). Case 3: A 4000-thread test with an initial timer tick of .01 second and an interval of 10,000 seconds (i.e. a timer that ticks only once) had very nearly the same performance in both cases: 6.3 seconds elapsed for the fixed kernel versus 5.5 seconds for the unfixed kernel. With fewer threads (eight in these tests), the Case 1 test ran in essentially the same time on both the modified and unmodified kernels (5.2 seconds versus 5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds versus 5.4 seconds but again with much better tick accuracy, .013 seconds per tick versus .025 seconds per tick for the unmodified kernel. Since the fix affected the rlimit code, I also tested soft and hard CPU limits. Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer running), the modified kernel was very slightly favored in that while it killed the process in 19.997 seconds of CPU time (5.002 seconds of wall time), only .003 seconds of that was system time, the rest was user time. The unmodified kernel killed the process in 20.001 seconds of CPU (5.014 seconds of wall time) of which .016 seconds was system time. Really, though, the results were too close to call. The results were essentially the same with no itimer running. Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds (where the hard limit would never be reached) and an itimer running, the modified kernel exhibited worse tick accuracy than the unmodified kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise, performance was almost indistinguishable. With no itimer running this test exhibited virtually identical behavior and times in both cases. In times past I did some limited performance testing. those results are below. On a four-cpu Opteron system without this fix, a sixteen-thread test executed in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On the same system with the fix, user and elapsed time were about the same, but system time dropped to 0.007 seconds. Performance with eight, four and one thread were comparable. Interestingly, the timer ticks with the fix seemed more accurate: The sixteen-thread test with the fix received 149543 ticks for 0.024 seconds per tick, while the same test without the fix received 58720 for 0.061 seconds per tick. Both cases were configured for an interval of 0.01 seconds. Again, the other tests were comparable. Each thread in this test computed the primes up to 25,000,000. I also did a test with a large number of threads, 100,000 threads, which is impossible without the fix. In this case each thread computed the primes only up to 10,000 (to make the runtime manageable). System time dominated, at 1546.968 seconds out of a total 2176.906 seconds (giving a user time of 629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite accurate. There is obviously no comparable test without the fix. Signed-off-by: Frank Mayhar <fmayhar@google.com> Cc: Roland McGrath <roland@redhat.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Cc: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-12 16:54:39 +00:00
/*
* Initialize POSIX timer handling for a thread group.
*/
static void posix_cpu_timers_init_group(struct signal_struct *sig)
{
unsigned long cpu_limit;
cpu_limit = READ_ONCE(sig->rlim[RLIMIT_CPU].rlim_cur);
if (cpu_limit != RLIM_INFINITY) {
sig->cputime_expires.prof_exp = cpu_limit * NSEC_PER_SEC;
sig->cputimer.running = true;
}
timers: fix itimer/many thread hang Overview This patch reworks the handling of POSIX CPU timers, including the ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together with the help of Roland McGrath, the owner and original writer of this code. The problem we ran into, and the reason for this rework, has to do with using a profiling timer in a process with a large number of threads. It appears that the performance of the old implementation of run_posix_cpu_timers() was at least O(n*3) (where "n" is the number of threads in a process) or worse. Everything is fine with an increasing number of threads until the time taken for that routine to run becomes the same as or greater than the tick time, at which point things degrade rather quickly. This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF." Code Changes This rework corrects the implementation of run_posix_cpu_timers() to make it run in constant time for a particular machine. (Performance may vary between one machine and another depending upon whether the kernel is built as single- or multiprocessor and, in the latter case, depending upon the number of running processors.) To do this, at each tick we now update fields in signal_struct as well as task_struct. The run_posix_cpu_timers() function uses those fields to make its decisions. We define a new structure, "task_cputime," to contain user, system and scheduler times and use these in appropriate places: struct task_cputime { cputime_t utime; cputime_t stime; unsigned long long sum_exec_runtime; }; This is included in the structure "thread_group_cputime," which is a new substructure of signal_struct and which varies for uniprocessor versus multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as a simple substructure, while for multiprocessor kernels it is a pointer: struct thread_group_cputime { struct task_cputime totals; }; struct thread_group_cputime { struct task_cputime *totals; }; We also add a new task_cputime substructure directly to signal_struct, to cache the earliest expiration of process-wide timers, and task_cputime also replaces the it_*_expires fields of task_struct (used for earliest expiration of thread timers). The "thread_group_cputime" structure contains process-wide timers that are updated via account_user_time() and friends. In the non-SMP case the structure is a simple aggregator; unfortunately in the SMP case that simplicity was not achievable due to cache-line contention between CPUs (in one measured case performance was actually _worse_ on a 16-cpu system than the same test on a 4-cpu system, due to this contention). For SMP, the thread_group_cputime counters are maintained as a per-cpu structure allocated using alloc_percpu(). The timer functions update only the timer field in the structure corresponding to the running CPU, obtained using per_cpu_ptr(). We define a set of inline functions in sched.h that we use to maintain the thread_group_cputime structure and hide the differences between UP and SMP implementations from the rest of the kernel. The thread_group_cputime_init() function initializes the thread_group_cputime structure for the given task. The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the out-of-line function thread_group_cputime_alloc_smp() to allocate and fill in the per-cpu structures and fields. The thread_group_cputime_free() function, also a no-op for UP, in SMP frees the per-cpu structures. The thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls thread_group_cputime_alloc() if the per-cpu structures haven't yet been allocated. The thread_group_cputime() function fills the task_cputime structure it is passed with the contents of the thread_group_cputime fields; in UP it's that simple but in SMP it must also safely check that tsk->signal is non-NULL (if it is it just uses the appropriate fields of task_struct) and, if so, sums the per-cpu values for each online CPU. Finally, the three functions account_group_user_time(), account_group_system_time() and account_group_exec_runtime() are used by timer functions to update the respective fields of the thread_group_cputime structure. Non-SMP operation is trivial and will not be mentioned further. The per-cpu structure is always allocated when a task creates its first new thread, via a call to thread_group_cputime_clone_thread() from copy_signal(). It is freed at process exit via a call to thread_group_cputime_free() from cleanup_signal(). All functions that formerly summed utime/stime/sum_sched_runtime values from from all threads in the thread group now use thread_group_cputime() to snapshot the values in the thread_group_cputime structure or the values in the task structure itself if the per-cpu structure hasn't been allocated. Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit. The run_posix_cpu_timers() function has been split into a fast path and a slow path; the former safely checks whether there are any expired thread timers and, if not, just returns, while the slow path does the heavy lifting. With the dedicated thread group fields, timers are no longer "rebalanced" and the process_timer_rebalance() function and related code has gone away. All summing loops are gone and all code that used them now uses the thread_group_cputime() inline. When process-wide timers are set, the new task_cputime structure in signal_struct is used to cache the earliest expiration; this is checked in the fast path. Performance The fix appears not to add significant overhead to existing operations. It generally performs the same as the current code except in two cases, one in which it performs slightly worse (Case 5 below) and one in which it performs very significantly better (Case 2 below). Overall it's a wash except in those two cases. I've since done somewhat more involved testing on a dual-core Opteron system. Case 1: With no itimer running, for a test with 100,000 threads, the fixed kernel took 1428.5 seconds, 513 seconds more than the unfixed system, all of which was spent in the system. There were twice as many voluntary context switches with the fix as without it. Case 2: With an itimer running at .01 second ticks and 4000 threads (the most an unmodified kernel can handle), the fixed kernel ran the test in eight percent of the time (5.8 seconds as opposed to 70 seconds) and had better tick accuracy (.012 seconds per tick as opposed to .023 seconds per tick). Case 3: A 4000-thread test with an initial timer tick of .01 second and an interval of 10,000 seconds (i.e. a timer that ticks only once) had very nearly the same performance in both cases: 6.3 seconds elapsed for the fixed kernel versus 5.5 seconds for the unfixed kernel. With fewer threads (eight in these tests), the Case 1 test ran in essentially the same time on both the modified and unmodified kernels (5.2 seconds versus 5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds versus 5.4 seconds but again with much better tick accuracy, .013 seconds per tick versus .025 seconds per tick for the unmodified kernel. Since the fix affected the rlimit code, I also tested soft and hard CPU limits. Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer running), the modified kernel was very slightly favored in that while it killed the process in 19.997 seconds of CPU time (5.002 seconds of wall time), only .003 seconds of that was system time, the rest was user time. The unmodified kernel killed the process in 20.001 seconds of CPU (5.014 seconds of wall time) of which .016 seconds was system time. Really, though, the results were too close to call. The results were essentially the same with no itimer running. Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds (where the hard limit would never be reached) and an itimer running, the modified kernel exhibited worse tick accuracy than the unmodified kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise, performance was almost indistinguishable. With no itimer running this test exhibited virtually identical behavior and times in both cases. In times past I did some limited performance testing. those results are below. On a four-cpu Opteron system without this fix, a sixteen-thread test executed in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On the same system with the fix, user and elapsed time were about the same, but system time dropped to 0.007 seconds. Performance with eight, four and one thread were comparable. Interestingly, the timer ticks with the fix seemed more accurate: The sixteen-thread test with the fix received 149543 ticks for 0.024 seconds per tick, while the same test without the fix received 58720 for 0.061 seconds per tick. Both cases were configured for an interval of 0.01 seconds. Again, the other tests were comparable. Each thread in this test computed the primes up to 25,000,000. I also did a test with a large number of threads, 100,000 threads, which is impossible without the fix. In this case each thread computed the primes only up to 10,000 (to make the runtime manageable). System time dominated, at 1546.968 seconds out of a total 2176.906 seconds (giving a user time of 629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite accurate. There is obviously no comparable test without the fix. Signed-off-by: Frank Mayhar <fmayhar@google.com> Cc: Roland McGrath <roland@redhat.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Cc: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-12 16:54:39 +00:00
/* The timer lists. */
INIT_LIST_HEAD(&sig->cpu_timers[0]);
INIT_LIST_HEAD(&sig->cpu_timers[1]);
INIT_LIST_HEAD(&sig->cpu_timers[2]);
}
#else
static inline void posix_cpu_timers_init_group(struct signal_struct *sig) { }
#endif
timers: fix itimer/many thread hang Overview This patch reworks the handling of POSIX CPU timers, including the ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together with the help of Roland McGrath, the owner and original writer of this code. The problem we ran into, and the reason for this rework, has to do with using a profiling timer in a process with a large number of threads. It appears that the performance of the old implementation of run_posix_cpu_timers() was at least O(n*3) (where "n" is the number of threads in a process) or worse. Everything is fine with an increasing number of threads until the time taken for that routine to run becomes the same as or greater than the tick time, at which point things degrade rather quickly. This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF." Code Changes This rework corrects the implementation of run_posix_cpu_timers() to make it run in constant time for a particular machine. (Performance may vary between one machine and another depending upon whether the kernel is built as single- or multiprocessor and, in the latter case, depending upon the number of running processors.) To do this, at each tick we now update fields in signal_struct as well as task_struct. The run_posix_cpu_timers() function uses those fields to make its decisions. We define a new structure, "task_cputime," to contain user, system and scheduler times and use these in appropriate places: struct task_cputime { cputime_t utime; cputime_t stime; unsigned long long sum_exec_runtime; }; This is included in the structure "thread_group_cputime," which is a new substructure of signal_struct and which varies for uniprocessor versus multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as a simple substructure, while for multiprocessor kernels it is a pointer: struct thread_group_cputime { struct task_cputime totals; }; struct thread_group_cputime { struct task_cputime *totals; }; We also add a new task_cputime substructure directly to signal_struct, to cache the earliest expiration of process-wide timers, and task_cputime also replaces the it_*_expires fields of task_struct (used for earliest expiration of thread timers). The "thread_group_cputime" structure contains process-wide timers that are updated via account_user_time() and friends. In the non-SMP case the structure is a simple aggregator; unfortunately in the SMP case that simplicity was not achievable due to cache-line contention between CPUs (in one measured case performance was actually _worse_ on a 16-cpu system than the same test on a 4-cpu system, due to this contention). For SMP, the thread_group_cputime counters are maintained as a per-cpu structure allocated using alloc_percpu(). The timer functions update only the timer field in the structure corresponding to the running CPU, obtained using per_cpu_ptr(). We define a set of inline functions in sched.h that we use to maintain the thread_group_cputime structure and hide the differences between UP and SMP implementations from the rest of the kernel. The thread_group_cputime_init() function initializes the thread_group_cputime structure for the given task. The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the out-of-line function thread_group_cputime_alloc_smp() to allocate and fill in the per-cpu structures and fields. The thread_group_cputime_free() function, also a no-op for UP, in SMP frees the per-cpu structures. The thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls thread_group_cputime_alloc() if the per-cpu structures haven't yet been allocated. The thread_group_cputime() function fills the task_cputime structure it is passed with the contents of the thread_group_cputime fields; in UP it's that simple but in SMP it must also safely check that tsk->signal is non-NULL (if it is it just uses the appropriate fields of task_struct) and, if so, sums the per-cpu values for each online CPU. Finally, the three functions account_group_user_time(), account_group_system_time() and account_group_exec_runtime() are used by timer functions to update the respective fields of the thread_group_cputime structure. Non-SMP operation is trivial and will not be mentioned further. The per-cpu structure is always allocated when a task creates its first new thread, via a call to thread_group_cputime_clone_thread() from copy_signal(). It is freed at process exit via a call to thread_group_cputime_free() from cleanup_signal(). All functions that formerly summed utime/stime/sum_sched_runtime values from from all threads in the thread group now use thread_group_cputime() to snapshot the values in the thread_group_cputime structure or the values in the task structure itself if the per-cpu structure hasn't been allocated. Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit. The run_posix_cpu_timers() function has been split into a fast path and a slow path; the former safely checks whether there are any expired thread timers and, if not, just returns, while the slow path does the heavy lifting. With the dedicated thread group fields, timers are no longer "rebalanced" and the process_timer_rebalance() function and related code has gone away. All summing loops are gone and all code that used them now uses the thread_group_cputime() inline. When process-wide timers are set, the new task_cputime structure in signal_struct is used to cache the earliest expiration; this is checked in the fast path. Performance The fix appears not to add significant overhead to existing operations. It generally performs the same as the current code except in two cases, one in which it performs slightly worse (Case 5 below) and one in which it performs very significantly better (Case 2 below). Overall it's a wash except in those two cases. I've since done somewhat more involved testing on a dual-core Opteron system. Case 1: With no itimer running, for a test with 100,000 threads, the fixed kernel took 1428.5 seconds, 513 seconds more than the unfixed system, all of which was spent in the system. There were twice as many voluntary context switches with the fix as without it. Case 2: With an itimer running at .01 second ticks and 4000 threads (the most an unmodified kernel can handle), the fixed kernel ran the test in eight percent of the time (5.8 seconds as opposed to 70 seconds) and had better tick accuracy (.012 seconds per tick as opposed to .023 seconds per tick). Case 3: A 4000-thread test with an initial timer tick of .01 second and an interval of 10,000 seconds (i.e. a timer that ticks only once) had very nearly the same performance in both cases: 6.3 seconds elapsed for the fixed kernel versus 5.5 seconds for the unfixed kernel. With fewer threads (eight in these tests), the Case 1 test ran in essentially the same time on both the modified and unmodified kernels (5.2 seconds versus 5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds versus 5.4 seconds but again with much better tick accuracy, .013 seconds per tick versus .025 seconds per tick for the unmodified kernel. Since the fix affected the rlimit code, I also tested soft and hard CPU limits. Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer running), the modified kernel was very slightly favored in that while it killed the process in 19.997 seconds of CPU time (5.002 seconds of wall time), only .003 seconds of that was system time, the rest was user time. The unmodified kernel killed the process in 20.001 seconds of CPU (5.014 seconds of wall time) of which .016 seconds was system time. Really, though, the results were too close to call. The results were essentially the same with no itimer running. Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds (where the hard limit would never be reached) and an itimer running, the modified kernel exhibited worse tick accuracy than the unmodified kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise, performance was almost indistinguishable. With no itimer running this test exhibited virtually identical behavior and times in both cases. In times past I did some limited performance testing. those results are below. On a four-cpu Opteron system without this fix, a sixteen-thread test executed in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On the same system with the fix, user and elapsed time were about the same, but system time dropped to 0.007 seconds. Performance with eight, four and one thread were comparable. Interestingly, the timer ticks with the fix seemed more accurate: The sixteen-thread test with the fix received 149543 ticks for 0.024 seconds per tick, while the same test without the fix received 58720 for 0.061 seconds per tick. Both cases were configured for an interval of 0.01 seconds. Again, the other tests were comparable. Each thread in this test computed the primes up to 25,000,000. I also did a test with a large number of threads, 100,000 threads, which is impossible without the fix. In this case each thread computed the primes only up to 10,000 (to make the runtime manageable). System time dominated, at 1546.968 seconds out of a total 2176.906 seconds (giving a user time of 629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite accurate. There is obviously no comparable test without the fix. Signed-off-by: Frank Mayhar <fmayhar@google.com> Cc: Roland McGrath <roland@redhat.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Cc: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-12 16:54:39 +00:00
static int copy_signal(unsigned long clone_flags, struct task_struct *tsk)
{
struct signal_struct *sig;
clone(): fix race between copy_process() and de_thread() Spotted by Hiroshi Shimamoto who also provided the test-case below. copy_process() uses signal->count as a reference counter, but it is not. This test case #include <sys/types.h> #include <sys/wait.h> #include <unistd.h> #include <stdio.h> #include <errno.h> #include <pthread.h> void *null_thread(void *p) { for (;;) sleep(1); return NULL; } void *exec_thread(void *p) { execl("/bin/true", "/bin/true", NULL); return null_thread(p); } int main(int argc, char **argv) { for (;;) { pid_t pid; int ret, status; pid = fork(); if (pid < 0) break; if (!pid) { pthread_t tid; pthread_create(&tid, NULL, exec_thread, NULL); for (;;) pthread_create(&tid, NULL, null_thread, NULL); } do { ret = waitpid(pid, &status, 0); } while (ret == -1 && errno == EINTR); } return 0; } quickly creates an unkillable task. If copy_process(CLONE_THREAD) races with de_thread() copy_signal()->atomic(signal->count) breaks the signal->notify_count logic, and the execing thread can hang forever in kernel space. Change copy_process() to increment count/live only when we know for sure we can't fail. In this case the forked thread will take care of its reference to signal correctly. If copy_process() fails, check CLONE_THREAD flag. If it it set - do nothing, the counters were not changed and current belongs to the same thread group. If it is not set, ->signal must be released in any case (and ->count must be == 1), the forked child is the only thread in the thread group. We need more cleanups here, in particular signal->count should not be used by de_thread/__exit_signal at all. This patch only fixes the bug. Reported-by: Hiroshi Shimamoto <h-shimamoto@ct.jp.nec.com> Tested-by: Hiroshi Shimamoto <h-shimamoto@ct.jp.nec.com> Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: <stable@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-08-26 21:29:24 +00:00
if (clone_flags & CLONE_THREAD)
return 0;
sig = kmem_cache_zalloc(signal_cachep, GFP_KERNEL);
tsk->signal = sig;
if (!sig)
return -ENOMEM;
sig->nr_threads = 1;
atomic_set(&sig->live, 1);
atomic_set(&sig->sigcnt, 1);
introduce for_each_thread() to replace the buggy while_each_thread() while_each_thread() and next_thread() should die, almost every lockless usage is wrong. 1. Unless g == current, the lockless while_each_thread() is not safe. while_each_thread(g, t) can loop forever if g exits, next_thread() can't reach the unhashed thread in this case. Note that this can happen even if g is the group leader, it can exec. 2. Even if while_each_thread() itself was correct, people often use it wrongly. It was never safe to just take rcu_read_lock() and loop unless you verify that pid_alive(g) == T, even the first next_thread() can point to the already freed/reused memory. This patch adds signal_struct->thread_head and task->thread_node to create the normal rcu-safe list with the stable head. The new for_each_thread(g, t) helper is always safe under rcu_read_lock() as long as this task_struct can't go away. Note: of course it is ugly to have both task_struct->thread_node and the old task_struct->thread_group, we will kill it later, after we change the users of while_each_thread() to use for_each_thread(). Perhaps we can kill it even before we convert all users, we can reimplement next_thread(t) using the new thread_head/thread_node. But we can't do this right now because this will lead to subtle behavioural changes. For example, do/while_each_thread() always sees at least one task, while for_each_thread() can do nothing if the whole thread group has died. Or thread_group_empty(), currently its semantics is not clear unless thread_group_leader(p) and we need to audit the callers before we can change it. So this patch adds the new interface which has to coexist with the old one for some time, hopefully the next changes will be more or less straightforward and the old one will go away soon. Signed-off-by: Oleg Nesterov <oleg@redhat.com> Reviewed-by: Sergey Dyasly <dserrg@gmail.com> Tested-by: Sergey Dyasly <dserrg@gmail.com> Reviewed-by: Sameer Nanda <snanda@chromium.org> Acked-by: David Rientjes <rientjes@google.com> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: Frederic Weisbecker <fweisbec@gmail.com> Cc: Mandeep Singh Baines <msb@chromium.org> Cc: "Ma, Xindong" <xindong.ma@intel.com> Cc: Michal Hocko <mhocko@suse.cz> Cc: "Tu, Xiaobing" <xiaobing.tu@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-01-21 23:49:56 +00:00
/* list_add(thread_node, thread_head) without INIT_LIST_HEAD() */
sig->thread_head = (struct list_head)LIST_HEAD_INIT(tsk->thread_node);
tsk->thread_node = (struct list_head)LIST_HEAD_INIT(sig->thread_head);
init_waitqueue_head(&sig->wait_chldexit);
sig->curr_target = tsk;
init_sigpending(&sig->shared_pending);
time, signal: Protect resource use statistics with seqlock Both times() and clock_gettime(CLOCK_PROCESS_CPUTIME_ID) have scalability issues on large systems, due to both functions being serialized with a lock. The lock protects against reporting a wrong value, due to a thread in the task group exiting, its statistics reporting up to the signal struct, and that exited task's statistics being counted twice (or not at all). Protecting that with a lock results in times() and clock_gettime() being completely serialized on large systems. This can be fixed by using a seqlock around the events that gather and propagate statistics. As an additional benefit, the protection code can be moved into thread_group_cputime(), slightly simplifying the calling functions. In the case of posix_cpu_clock_get_task() things can be simplified a lot, because the calling function already ensures that the task sticks around, and the rest is now taken care of in thread_group_cputime(). This way the statistics reporting code can run lockless. Signed-off-by: Rik van Riel <riel@redhat.com> Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org> Cc: Alex Thorlton <athorlton@sgi.com> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Daeseok Youn <daeseok.youn@gmail.com> Cc: David Rientjes <rientjes@google.com> Cc: Dongsheng Yang <yangds.fnst@cn.fujitsu.com> Cc: Geert Uytterhoeven <geert@linux-m68k.org> Cc: Guillaume Morin <guillaume@morinfr.org> Cc: Ionut Alexa <ionut.m.alexa@gmail.com> Cc: Kees Cook <keescook@chromium.org> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Li Zefan <lizefan@huawei.com> Cc: Michal Hocko <mhocko@suse.cz> Cc: Michal Schmidt <mschmidt@redhat.com> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Vladimir Davydov <vdavydov@parallels.com> Cc: umgwanakikbuti@gmail.com Cc: fweisbec@gmail.com Cc: srao@redhat.com Cc: lwoodman@redhat.com Cc: atheurer@redhat.com Link: http://lkml.kernel.org/r/20140816134010.26a9b572@annuminas.surriel.com Signed-off-by: Ingo Molnar <mingo@kernel.org>
2014-08-16 17:40:10 +00:00
seqlock_init(&sig->stats_lock);
prev_cputime_init(&sig->prev_cputime);
#ifdef CONFIG_POSIX_TIMERS
INIT_LIST_HEAD(&sig->posix_timers);
hrtimer_init(&sig->real_timer, CLOCK_MONOTONIC, HRTIMER_MODE_REL);
sig->real_timer.function = it_real_fn;
#endif
task_lock(current->group_leader);
memcpy(sig->rlim, current->signal->rlim, sizeof sig->rlim);
task_unlock(current->group_leader);
posix_cpu_timers_init_group(sig);
Audit: add TTY input auditing Add TTY input auditing, used to audit system administrator's actions. This is required by various security standards such as DCID 6/3 and PCI to provide non-repudiation of administrator's actions and to allow a review of past actions if the administrator seems to overstep their duties or if the system becomes misconfigured for unknown reasons. These requirements do not make it necessary to audit TTY output as well. Compared to an user-space keylogger, this approach records TTY input using the audit subsystem, correlated with other audit events, and it is completely transparent to the user-space application (e.g. the console ioctls still work). TTY input auditing works on a higher level than auditing all system calls within the session, which would produce an overwhelming amount of mostly useless audit events. Add an "audit_tty" attribute, inherited across fork (). Data read from TTYs by process with the attribute is sent to the audit subsystem by the kernel. The audit netlink interface is extended to allow modifying the audit_tty attribute, and to allow sending explanatory audit events from user-space (for example, a shell might send an event containing the final command, after the interactive command-line editing and history expansion is performed, which might be difficult to decipher from the TTY input alone). Because the "audit_tty" attribute is inherited across fork (), it would be set e.g. for sshd restarted within an audited session. To prevent this, the audit_tty attribute is cleared when a process with no open TTY file descriptors (e.g. after daemon startup) opens a TTY. See https://www.redhat.com/archives/linux-audit/2007-June/msg00000.html for a more detailed rationale document for an older version of this patch. [akpm@linux-foundation.org: build fix] Signed-off-by: Miloslav Trmac <mitr@redhat.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Alan Cox <alan@lxorguk.ukuu.org.uk> Cc: Paul Fulghum <paulkf@microgate.com> Cc: Casey Schaufler <casey@schaufler-ca.com> Cc: Steve Grubb <sgrubb@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-07-16 06:40:56 +00:00
tty_audit_fork(sig);
sched: Add 'autogroup' scheduling feature: automated per session task groups A recurring complaint from CFS users is that parallel kbuild has a negative impact on desktop interactivity. This patch implements an idea from Linus, to automatically create task groups. Currently, only per session autogroups are implemented, but the patch leaves the way open for enhancement. Implementation: each task's signal struct contains an inherited pointer to a refcounted autogroup struct containing a task group pointer, the default for all tasks pointing to the init_task_group. When a task calls setsid(), a new task group is created, the process is moved into the new task group, and a reference to the preveious task group is dropped. Child processes inherit this task group thereafter, and increase it's refcount. When the last thread of a process exits, the process's reference is dropped, such that when the last process referencing an autogroup exits, the autogroup is destroyed. At runqueue selection time, IFF a task has no cgroup assignment, its current autogroup is used. Autogroup bandwidth is controllable via setting it's nice level through the proc filesystem: cat /proc/<pid>/autogroup Displays the task's group and the group's nice level. echo <nice level> > /proc/<pid>/autogroup Sets the task group's shares to the weight of nice <level> task. Setting nice level is rate limited for !admin users due to the abuse risk of task group locking. The feature is enabled from boot by default if CONFIG_SCHED_AUTOGROUP=y is selected, but can be disabled via the boot option noautogroup, and can also be turned on/off on the fly via: echo [01] > /proc/sys/kernel/sched_autogroup_enabled ... which will automatically move tasks to/from the root task group. Signed-off-by: Mike Galbraith <efault@gmx.de> Acked-by: Linus Torvalds <torvalds@linux-foundation.org> Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: Markus Trippelsdorf <markus@trippelsdorf.de> Cc: Mathieu Desnoyers <mathieu.desnoyers@efficios.com> Cc: Paul Turner <pjt@google.com> Cc: Oleg Nesterov <oleg@redhat.com> [ Removed the task_group_path() debug code, and fixed !EVENTFD build failure. ] Signed-off-by: Ingo Molnar <mingo@elte.hu> LKML-Reference: <1290281700.28711.9.camel@maggy.simson.net> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2010-11-30 13:18:03 +00:00
sched_autogroup_fork(sig);
Audit: add TTY input auditing Add TTY input auditing, used to audit system administrator's actions. This is required by various security standards such as DCID 6/3 and PCI to provide non-repudiation of administrator's actions and to allow a review of past actions if the administrator seems to overstep their duties or if the system becomes misconfigured for unknown reasons. These requirements do not make it necessary to audit TTY output as well. Compared to an user-space keylogger, this approach records TTY input using the audit subsystem, correlated with other audit events, and it is completely transparent to the user-space application (e.g. the console ioctls still work). TTY input auditing works on a higher level than auditing all system calls within the session, which would produce an overwhelming amount of mostly useless audit events. Add an "audit_tty" attribute, inherited across fork (). Data read from TTYs by process with the attribute is sent to the audit subsystem by the kernel. The audit netlink interface is extended to allow modifying the audit_tty attribute, and to allow sending explanatory audit events from user-space (for example, a shell might send an event containing the final command, after the interactive command-line editing and history expansion is performed, which might be difficult to decipher from the TTY input alone). Because the "audit_tty" attribute is inherited across fork (), it would be set e.g. for sshd restarted within an audited session. To prevent this, the audit_tty attribute is cleared when a process with no open TTY file descriptors (e.g. after daemon startup) opens a TTY. See https://www.redhat.com/archives/linux-audit/2007-June/msg00000.html for a more detailed rationale document for an older version of this patch. [akpm@linux-foundation.org: build fix] Signed-off-by: Miloslav Trmac <mitr@redhat.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Alan Cox <alan@lxorguk.ukuu.org.uk> Cc: Paul Fulghum <paulkf@microgate.com> Cc: Casey Schaufler <casey@schaufler-ca.com> Cc: Steve Grubb <sgrubb@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-07-16 06:40:56 +00:00
oom: badness heuristic rewrite This a complete rewrite of the oom killer's badness() heuristic which is used to determine which task to kill in oom conditions. The goal is to make it as simple and predictable as possible so the results are better understood and we end up killing the task which will lead to the most memory freeing while still respecting the fine-tuning from userspace. Instead of basing the heuristic on mm->total_vm for each task, the task's rss and swap space is used instead. This is a better indication of the amount of memory that will be freeable if the oom killed task is chosen and subsequently exits. This helps specifically in cases where KDE or GNOME is chosen for oom kill on desktop systems instead of a memory hogging task. The baseline for the heuristic is a proportion of memory that each task is currently using in memory plus swap compared to the amount of "allowable" memory. "Allowable," in this sense, means the system-wide resources for unconstrained oom conditions, the set of mempolicy nodes, the mems attached to current's cpuset, or a memory controller's limit. The proportion is given on a scale of 0 (never kill) to 1000 (always kill), roughly meaning that if a task has a badness() score of 500 that the task consumes approximately 50% of allowable memory resident in RAM or in swap space. The proportion is always relative to the amount of "allowable" memory and not the total amount of RAM systemwide so that mempolicies and cpusets may operate in isolation; they shall not need to know the true size of the machine on which they are running if they are bound to a specific set of nodes or mems, respectively. Root tasks are given 3% extra memory just like __vm_enough_memory() provides in LSMs. In the event of two tasks consuming similar amounts of memory, it is generally better to save root's task. Because of the change in the badness() heuristic's baseline, it is also necessary to introduce a new user interface to tune it. It's not possible to redefine the meaning of /proc/pid/oom_adj with a new scale since the ABI cannot be changed for backward compatability. Instead, a new tunable, /proc/pid/oom_score_adj, is added that ranges from -1000 to +1000. It may be used to polarize the heuristic such that certain tasks are never considered for oom kill while others may always be considered. The value is added directly into the badness() score so a value of -500, for example, means to discount 50% of its memory consumption in comparison to other tasks either on the system, bound to the mempolicy, in the cpuset, or sharing the same memory controller. /proc/pid/oom_adj is changed so that its meaning is rescaled into the units used by /proc/pid/oom_score_adj, and vice versa. Changing one of these per-task tunables will rescale the value of the other to an equivalent meaning. Although /proc/pid/oom_adj was originally defined as a bitshift on the badness score, it now shares the same linear growth as /proc/pid/oom_score_adj but with different granularity. This is required so the ABI is not broken with userspace applications and allows oom_adj to be deprecated for future removal. Signed-off-by: David Rientjes <rientjes@google.com> Cc: Nick Piggin <npiggin@suse.de> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Balbir Singh <balbir@in.ibm.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-08-10 00:19:46 +00:00
sig->oom_score_adj = current->signal->oom_score_adj;
sig->oom_score_adj_min = current->signal->oom_score_adj_min;
oom: move oom_adj value from task_struct to signal_struct Currently, OOM logic callflow is here. __out_of_memory() select_bad_process() for each task badness() calculate badness of one task oom_kill_process() search child oom_kill_task() kill target task and mm shared tasks with it example, process-A have two thread, thread-A and thread-B and it have very fat memory and each thread have following oom_adj and oom_score. thread-A: oom_adj = OOM_DISABLE, oom_score = 0 thread-B: oom_adj = 0, oom_score = very-high Then, select_bad_process() select thread-B, but oom_kill_task() refuse kill the task because thread-A have OOM_DISABLE. Thus __out_of_memory() call select_bad_process() again. but select_bad_process() select the same task. It mean kernel fall in livelock. The fact is, select_bad_process() must select killable task. otherwise OOM logic go into livelock. And root cause is, oom_adj shouldn't be per-thread value. it should be per-process value because OOM-killer kill a process, not thread. Thus This patch moves oomkilladj (now more appropriately named oom_adj) from struct task_struct to struct signal_struct. it naturally prevent select_bad_process() choose wrong task. Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Paul Menage <menage@google.com> Cc: David Rientjes <rientjes@google.com> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Rik van Riel <riel@redhat.com> Cc: Oleg Nesterov <oleg@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-22 00:03:13 +00:00
mutex_init(&sig->cred_guard_mutex);
return 0;
}
static void copy_seccomp(struct task_struct *p)
{
#ifdef CONFIG_SECCOMP
/*
* Must be called with sighand->lock held, which is common to
* all threads in the group. Holding cred_guard_mutex is not
* needed because this new task is not yet running and cannot
* be racing exec.
*/
assert_spin_locked(&current->sighand->siglock);
/* Ref-count the new filter user, and assign it. */
get_seccomp_filter(current);
p->seccomp = current->seccomp;
/*
* Explicitly enable no_new_privs here in case it got set
* between the task_struct being duplicated and holding the
* sighand lock. The seccomp state and nnp must be in sync.
*/
if (task_no_new_privs(current))
task_set_no_new_privs(p);
/*
* If the parent gained a seccomp mode after copying thread
* flags and between before we held the sighand lock, we have
* to manually enable the seccomp thread flag here.
*/
if (p->seccomp.mode != SECCOMP_MODE_DISABLED)
set_tsk_thread_flag(p, TIF_SECCOMP);
#endif
}
SYSCALL_DEFINE1(set_tid_address, int __user *, tidptr)
{
current->clear_child_tid = tidptr;
return task_pid_vnr(current);
}
static void rt_mutex_init_task(struct task_struct *p)
{
raw_spin_lock_init(&p->pi_lock);
#ifdef CONFIG_RT_MUTEXES
p->pi_waiters = RB_ROOT;
p->pi_waiters_leftmost = NULL;
p->pi_blocked_on = NULL;
#endif
}
#ifdef CONFIG_POSIX_TIMERS
timers: fix itimer/many thread hang Overview This patch reworks the handling of POSIX CPU timers, including the ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together with the help of Roland McGrath, the owner and original writer of this code. The problem we ran into, and the reason for this rework, has to do with using a profiling timer in a process with a large number of threads. It appears that the performance of the old implementation of run_posix_cpu_timers() was at least O(n*3) (where "n" is the number of threads in a process) or worse. Everything is fine with an increasing number of threads until the time taken for that routine to run becomes the same as or greater than the tick time, at which point things degrade rather quickly. This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF." Code Changes This rework corrects the implementation of run_posix_cpu_timers() to make it run in constant time for a particular machine. (Performance may vary between one machine and another depending upon whether the kernel is built as single- or multiprocessor and, in the latter case, depending upon the number of running processors.) To do this, at each tick we now update fields in signal_struct as well as task_struct. The run_posix_cpu_timers() function uses those fields to make its decisions. We define a new structure, "task_cputime," to contain user, system and scheduler times and use these in appropriate places: struct task_cputime { cputime_t utime; cputime_t stime; unsigned long long sum_exec_runtime; }; This is included in the structure "thread_group_cputime," which is a new substructure of signal_struct and which varies for uniprocessor versus multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as a simple substructure, while for multiprocessor kernels it is a pointer: struct thread_group_cputime { struct task_cputime totals; }; struct thread_group_cputime { struct task_cputime *totals; }; We also add a new task_cputime substructure directly to signal_struct, to cache the earliest expiration of process-wide timers, and task_cputime also replaces the it_*_expires fields of task_struct (used for earliest expiration of thread timers). The "thread_group_cputime" structure contains process-wide timers that are updated via account_user_time() and friends. In the non-SMP case the structure is a simple aggregator; unfortunately in the SMP case that simplicity was not achievable due to cache-line contention between CPUs (in one measured case performance was actually _worse_ on a 16-cpu system than the same test on a 4-cpu system, due to this contention). For SMP, the thread_group_cputime counters are maintained as a per-cpu structure allocated using alloc_percpu(). The timer functions update only the timer field in the structure corresponding to the running CPU, obtained using per_cpu_ptr(). We define a set of inline functions in sched.h that we use to maintain the thread_group_cputime structure and hide the differences between UP and SMP implementations from the rest of the kernel. The thread_group_cputime_init() function initializes the thread_group_cputime structure for the given task. The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the out-of-line function thread_group_cputime_alloc_smp() to allocate and fill in the per-cpu structures and fields. The thread_group_cputime_free() function, also a no-op for UP, in SMP frees the per-cpu structures. The thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls thread_group_cputime_alloc() if the per-cpu structures haven't yet been allocated. The thread_group_cputime() function fills the task_cputime structure it is passed with the contents of the thread_group_cputime fields; in UP it's that simple but in SMP it must also safely check that tsk->signal is non-NULL (if it is it just uses the appropriate fields of task_struct) and, if so, sums the per-cpu values for each online CPU. Finally, the three functions account_group_user_time(), account_group_system_time() and account_group_exec_runtime() are used by timer functions to update the respective fields of the thread_group_cputime structure. Non-SMP operation is trivial and will not be mentioned further. The per-cpu structure is always allocated when a task creates its first new thread, via a call to thread_group_cputime_clone_thread() from copy_signal(). It is freed at process exit via a call to thread_group_cputime_free() from cleanup_signal(). All functions that formerly summed utime/stime/sum_sched_runtime values from from all threads in the thread group now use thread_group_cputime() to snapshot the values in the thread_group_cputime structure or the values in the task structure itself if the per-cpu structure hasn't been allocated. Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit. The run_posix_cpu_timers() function has been split into a fast path and a slow path; the former safely checks whether there are any expired thread timers and, if not, just returns, while the slow path does the heavy lifting. With the dedicated thread group fields, timers are no longer "rebalanced" and the process_timer_rebalance() function and related code has gone away. All summing loops are gone and all code that used them now uses the thread_group_cputime() inline. When process-wide timers are set, the new task_cputime structure in signal_struct is used to cache the earliest expiration; this is checked in the fast path. Performance The fix appears not to add significant overhead to existing operations. It generally performs the same as the current code except in two cases, one in which it performs slightly worse (Case 5 below) and one in which it performs very significantly better (Case 2 below). Overall it's a wash except in those two cases. I've since done somewhat more involved testing on a dual-core Opteron system. Case 1: With no itimer running, for a test with 100,000 threads, the fixed kernel took 1428.5 seconds, 513 seconds more than the unfixed system, all of which was spent in the system. There were twice as many voluntary context switches with the fix as without it. Case 2: With an itimer running at .01 second ticks and 4000 threads (the most an unmodified kernel can handle), the fixed kernel ran the test in eight percent of the time (5.8 seconds as opposed to 70 seconds) and had better tick accuracy (.012 seconds per tick as opposed to .023 seconds per tick). Case 3: A 4000-thread test with an initial timer tick of .01 second and an interval of 10,000 seconds (i.e. a timer that ticks only once) had very nearly the same performance in both cases: 6.3 seconds elapsed for the fixed kernel versus 5.5 seconds for the unfixed kernel. With fewer threads (eight in these tests), the Case 1 test ran in essentially the same time on both the modified and unmodified kernels (5.2 seconds versus 5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds versus 5.4 seconds but again with much better tick accuracy, .013 seconds per tick versus .025 seconds per tick for the unmodified kernel. Since the fix affected the rlimit code, I also tested soft and hard CPU limits. Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer running), the modified kernel was very slightly favored in that while it killed the process in 19.997 seconds of CPU time (5.002 seconds of wall time), only .003 seconds of that was system time, the rest was user time. The unmodified kernel killed the process in 20.001 seconds of CPU (5.014 seconds of wall time) of which .016 seconds was system time. Really, though, the results were too close to call. The results were essentially the same with no itimer running. Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds (where the hard limit would never be reached) and an itimer running, the modified kernel exhibited worse tick accuracy than the unmodified kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise, performance was almost indistinguishable. With no itimer running this test exhibited virtually identical behavior and times in both cases. In times past I did some limited performance testing. those results are below. On a four-cpu Opteron system without this fix, a sixteen-thread test executed in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On the same system with the fix, user and elapsed time were about the same, but system time dropped to 0.007 seconds. Performance with eight, four and one thread were comparable. Interestingly, the timer ticks with the fix seemed more accurate: The sixteen-thread test with the fix received 149543 ticks for 0.024 seconds per tick, while the same test without the fix received 58720 for 0.061 seconds per tick. Both cases were configured for an interval of 0.01 seconds. Again, the other tests were comparable. Each thread in this test computed the primes up to 25,000,000. I also did a test with a large number of threads, 100,000 threads, which is impossible without the fix. In this case each thread computed the primes only up to 10,000 (to make the runtime manageable). System time dominated, at 1546.968 seconds out of a total 2176.906 seconds (giving a user time of 629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite accurate. There is obviously no comparable test without the fix. Signed-off-by: Frank Mayhar <fmayhar@google.com> Cc: Roland McGrath <roland@redhat.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Cc: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-12 16:54:39 +00:00
/*
* Initialize POSIX timer handling for a single task.
*/
static void posix_cpu_timers_init(struct task_struct *tsk)
{
tsk->cputime_expires.prof_exp = 0;
tsk->cputime_expires.virt_exp = 0;
timers: fix itimer/many thread hang Overview This patch reworks the handling of POSIX CPU timers, including the ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together with the help of Roland McGrath, the owner and original writer of this code. The problem we ran into, and the reason for this rework, has to do with using a profiling timer in a process with a large number of threads. It appears that the performance of the old implementation of run_posix_cpu_timers() was at least O(n*3) (where "n" is the number of threads in a process) or worse. Everything is fine with an increasing number of threads until the time taken for that routine to run becomes the same as or greater than the tick time, at which point things degrade rather quickly. This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF." Code Changes This rework corrects the implementation of run_posix_cpu_timers() to make it run in constant time for a particular machine. (Performance may vary between one machine and another depending upon whether the kernel is built as single- or multiprocessor and, in the latter case, depending upon the number of running processors.) To do this, at each tick we now update fields in signal_struct as well as task_struct. The run_posix_cpu_timers() function uses those fields to make its decisions. We define a new structure, "task_cputime," to contain user, system and scheduler times and use these in appropriate places: struct task_cputime { cputime_t utime; cputime_t stime; unsigned long long sum_exec_runtime; }; This is included in the structure "thread_group_cputime," which is a new substructure of signal_struct and which varies for uniprocessor versus multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as a simple substructure, while for multiprocessor kernels it is a pointer: struct thread_group_cputime { struct task_cputime totals; }; struct thread_group_cputime { struct task_cputime *totals; }; We also add a new task_cputime substructure directly to signal_struct, to cache the earliest expiration of process-wide timers, and task_cputime also replaces the it_*_expires fields of task_struct (used for earliest expiration of thread timers). The "thread_group_cputime" structure contains process-wide timers that are updated via account_user_time() and friends. In the non-SMP case the structure is a simple aggregator; unfortunately in the SMP case that simplicity was not achievable due to cache-line contention between CPUs (in one measured case performance was actually _worse_ on a 16-cpu system than the same test on a 4-cpu system, due to this contention). For SMP, the thread_group_cputime counters are maintained as a per-cpu structure allocated using alloc_percpu(). The timer functions update only the timer field in the structure corresponding to the running CPU, obtained using per_cpu_ptr(). We define a set of inline functions in sched.h that we use to maintain the thread_group_cputime structure and hide the differences between UP and SMP implementations from the rest of the kernel. The thread_group_cputime_init() function initializes the thread_group_cputime structure for the given task. The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the out-of-line function thread_group_cputime_alloc_smp() to allocate and fill in the per-cpu structures and fields. The thread_group_cputime_free() function, also a no-op for UP, in SMP frees the per-cpu structures. The thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls thread_group_cputime_alloc() if the per-cpu structures haven't yet been allocated. The thread_group_cputime() function fills the task_cputime structure it is passed with the contents of the thread_group_cputime fields; in UP it's that simple but in SMP it must also safely check that tsk->signal is non-NULL (if it is it just uses the appropriate fields of task_struct) and, if so, sums the per-cpu values for each online CPU. Finally, the three functions account_group_user_time(), account_group_system_time() and account_group_exec_runtime() are used by timer functions to update the respective fields of the thread_group_cputime structure. Non-SMP operation is trivial and will not be mentioned further. The per-cpu structure is always allocated when a task creates its first new thread, via a call to thread_group_cputime_clone_thread() from copy_signal(). It is freed at process exit via a call to thread_group_cputime_free() from cleanup_signal(). All functions that formerly summed utime/stime/sum_sched_runtime values from from all threads in the thread group now use thread_group_cputime() to snapshot the values in the thread_group_cputime structure or the values in the task structure itself if the per-cpu structure hasn't been allocated. Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit. The run_posix_cpu_timers() function has been split into a fast path and a slow path; the former safely checks whether there are any expired thread timers and, if not, just returns, while the slow path does the heavy lifting. With the dedicated thread group fields, timers are no longer "rebalanced" and the process_timer_rebalance() function and related code has gone away. All summing loops are gone and all code that used them now uses the thread_group_cputime() inline. When process-wide timers are set, the new task_cputime structure in signal_struct is used to cache the earliest expiration; this is checked in the fast path. Performance The fix appears not to add significant overhead to existing operations. It generally performs the same as the current code except in two cases, one in which it performs slightly worse (Case 5 below) and one in which it performs very significantly better (Case 2 below). Overall it's a wash except in those two cases. I've since done somewhat more involved testing on a dual-core Opteron system. Case 1: With no itimer running, for a test with 100,000 threads, the fixed kernel took 1428.5 seconds, 513 seconds more than the unfixed system, all of which was spent in the system. There were twice as many voluntary context switches with the fix as without it. Case 2: With an itimer running at .01 second ticks and 4000 threads (the most an unmodified kernel can handle), the fixed kernel ran the test in eight percent of the time (5.8 seconds as opposed to 70 seconds) and had better tick accuracy (.012 seconds per tick as opposed to .023 seconds per tick). Case 3: A 4000-thread test with an initial timer tick of .01 second and an interval of 10,000 seconds (i.e. a timer that ticks only once) had very nearly the same performance in both cases: 6.3 seconds elapsed for the fixed kernel versus 5.5 seconds for the unfixed kernel. With fewer threads (eight in these tests), the Case 1 test ran in essentially the same time on both the modified and unmodified kernels (5.2 seconds versus 5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds versus 5.4 seconds but again with much better tick accuracy, .013 seconds per tick versus .025 seconds per tick for the unmodified kernel. Since the fix affected the rlimit code, I also tested soft and hard CPU limits. Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer running), the modified kernel was very slightly favored in that while it killed the process in 19.997 seconds of CPU time (5.002 seconds of wall time), only .003 seconds of that was system time, the rest was user time. The unmodified kernel killed the process in 20.001 seconds of CPU (5.014 seconds of wall time) of which .016 seconds was system time. Really, though, the results were too close to call. The results were essentially the same with no itimer running. Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds (where the hard limit would never be reached) and an itimer running, the modified kernel exhibited worse tick accuracy than the unmodified kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise, performance was almost indistinguishable. With no itimer running this test exhibited virtually identical behavior and times in both cases. In times past I did some limited performance testing. those results are below. On a four-cpu Opteron system without this fix, a sixteen-thread test executed in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On the same system with the fix, user and elapsed time were about the same, but system time dropped to 0.007 seconds. Performance with eight, four and one thread were comparable. Interestingly, the timer ticks with the fix seemed more accurate: The sixteen-thread test with the fix received 149543 ticks for 0.024 seconds per tick, while the same test without the fix received 58720 for 0.061 seconds per tick. Both cases were configured for an interval of 0.01 seconds. Again, the other tests were comparable. Each thread in this test computed the primes up to 25,000,000. I also did a test with a large number of threads, 100,000 threads, which is impossible without the fix. In this case each thread computed the primes only up to 10,000 (to make the runtime manageable). System time dominated, at 1546.968 seconds out of a total 2176.906 seconds (giving a user time of 629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite accurate. There is obviously no comparable test without the fix. Signed-off-by: Frank Mayhar <fmayhar@google.com> Cc: Roland McGrath <roland@redhat.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Cc: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-12 16:54:39 +00:00
tsk->cputime_expires.sched_exp = 0;
INIT_LIST_HEAD(&tsk->cpu_timers[0]);
INIT_LIST_HEAD(&tsk->cpu_timers[1]);
INIT_LIST_HEAD(&tsk->cpu_timers[2]);
}
#else
static inline void posix_cpu_timers_init(struct task_struct *tsk) { }
#endif
timers: fix itimer/many thread hang Overview This patch reworks the handling of POSIX CPU timers, including the ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together with the help of Roland McGrath, the owner and original writer of this code. The problem we ran into, and the reason for this rework, has to do with using a profiling timer in a process with a large number of threads. It appears that the performance of the old implementation of run_posix_cpu_timers() was at least O(n*3) (where "n" is the number of threads in a process) or worse. Everything is fine with an increasing number of threads until the time taken for that routine to run becomes the same as or greater than the tick time, at which point things degrade rather quickly. This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF." Code Changes This rework corrects the implementation of run_posix_cpu_timers() to make it run in constant time for a particular machine. (Performance may vary between one machine and another depending upon whether the kernel is built as single- or multiprocessor and, in the latter case, depending upon the number of running processors.) To do this, at each tick we now update fields in signal_struct as well as task_struct. The run_posix_cpu_timers() function uses those fields to make its decisions. We define a new structure, "task_cputime," to contain user, system and scheduler times and use these in appropriate places: struct task_cputime { cputime_t utime; cputime_t stime; unsigned long long sum_exec_runtime; }; This is included in the structure "thread_group_cputime," which is a new substructure of signal_struct and which varies for uniprocessor versus multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as a simple substructure, while for multiprocessor kernels it is a pointer: struct thread_group_cputime { struct task_cputime totals; }; struct thread_group_cputime { struct task_cputime *totals; }; We also add a new task_cputime substructure directly to signal_struct, to cache the earliest expiration of process-wide timers, and task_cputime also replaces the it_*_expires fields of task_struct (used for earliest expiration of thread timers). The "thread_group_cputime" structure contains process-wide timers that are updated via account_user_time() and friends. In the non-SMP case the structure is a simple aggregator; unfortunately in the SMP case that simplicity was not achievable due to cache-line contention between CPUs (in one measured case performance was actually _worse_ on a 16-cpu system than the same test on a 4-cpu system, due to this contention). For SMP, the thread_group_cputime counters are maintained as a per-cpu structure allocated using alloc_percpu(). The timer functions update only the timer field in the structure corresponding to the running CPU, obtained using per_cpu_ptr(). We define a set of inline functions in sched.h that we use to maintain the thread_group_cputime structure and hide the differences between UP and SMP implementations from the rest of the kernel. The thread_group_cputime_init() function initializes the thread_group_cputime structure for the given task. The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the out-of-line function thread_group_cputime_alloc_smp() to allocate and fill in the per-cpu structures and fields. The thread_group_cputime_free() function, also a no-op for UP, in SMP frees the per-cpu structures. The thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls thread_group_cputime_alloc() if the per-cpu structures haven't yet been allocated. The thread_group_cputime() function fills the task_cputime structure it is passed with the contents of the thread_group_cputime fields; in UP it's that simple but in SMP it must also safely check that tsk->signal is non-NULL (if it is it just uses the appropriate fields of task_struct) and, if so, sums the per-cpu values for each online CPU. Finally, the three functions account_group_user_time(), account_group_system_time() and account_group_exec_runtime() are used by timer functions to update the respective fields of the thread_group_cputime structure. Non-SMP operation is trivial and will not be mentioned further. The per-cpu structure is always allocated when a task creates its first new thread, via a call to thread_group_cputime_clone_thread() from copy_signal(). It is freed at process exit via a call to thread_group_cputime_free() from cleanup_signal(). All functions that formerly summed utime/stime/sum_sched_runtime values from from all threads in the thread group now use thread_group_cputime() to snapshot the values in the thread_group_cputime structure or the values in the task structure itself if the per-cpu structure hasn't been allocated. Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit. The run_posix_cpu_timers() function has been split into a fast path and a slow path; the former safely checks whether there are any expired thread timers and, if not, just returns, while the slow path does the heavy lifting. With the dedicated thread group fields, timers are no longer "rebalanced" and the process_timer_rebalance() function and related code has gone away. All summing loops are gone and all code that used them now uses the thread_group_cputime() inline. When process-wide timers are set, the new task_cputime structure in signal_struct is used to cache the earliest expiration; this is checked in the fast path. Performance The fix appears not to add significant overhead to existing operations. It generally performs the same as the current code except in two cases, one in which it performs slightly worse (Case 5 below) and one in which it performs very significantly better (Case 2 below). Overall it's a wash except in those two cases. I've since done somewhat more involved testing on a dual-core Opteron system. Case 1: With no itimer running, for a test with 100,000 threads, the fixed kernel took 1428.5 seconds, 513 seconds more than the unfixed system, all of which was spent in the system. There were twice as many voluntary context switches with the fix as without it. Case 2: With an itimer running at .01 second ticks and 4000 threads (the most an unmodified kernel can handle), the fixed kernel ran the test in eight percent of the time (5.8 seconds as opposed to 70 seconds) and had better tick accuracy (.012 seconds per tick as opposed to .023 seconds per tick). Case 3: A 4000-thread test with an initial timer tick of .01 second and an interval of 10,000 seconds (i.e. a timer that ticks only once) had very nearly the same performance in both cases: 6.3 seconds elapsed for the fixed kernel versus 5.5 seconds for the unfixed kernel. With fewer threads (eight in these tests), the Case 1 test ran in essentially the same time on both the modified and unmodified kernels (5.2 seconds versus 5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds versus 5.4 seconds but again with much better tick accuracy, .013 seconds per tick versus .025 seconds per tick for the unmodified kernel. Since the fix affected the rlimit code, I also tested soft and hard CPU limits. Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer running), the modified kernel was very slightly favored in that while it killed the process in 19.997 seconds of CPU time (5.002 seconds of wall time), only .003 seconds of that was system time, the rest was user time. The unmodified kernel killed the process in 20.001 seconds of CPU (5.014 seconds of wall time) of which .016 seconds was system time. Really, though, the results were too close to call. The results were essentially the same with no itimer running. Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds (where the hard limit would never be reached) and an itimer running, the modified kernel exhibited worse tick accuracy than the unmodified kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise, performance was almost indistinguishable. With no itimer running this test exhibited virtually identical behavior and times in both cases. In times past I did some limited performance testing. those results are below. On a four-cpu Opteron system without this fix, a sixteen-thread test executed in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On the same system with the fix, user and elapsed time were about the same, but system time dropped to 0.007 seconds. Performance with eight, four and one thread were comparable. Interestingly, the timer ticks with the fix seemed more accurate: The sixteen-thread test with the fix received 149543 ticks for 0.024 seconds per tick, while the same test without the fix received 58720 for 0.061 seconds per tick. Both cases were configured for an interval of 0.01 seconds. Again, the other tests were comparable. Each thread in this test computed the primes up to 25,000,000. I also did a test with a large number of threads, 100,000 threads, which is impossible without the fix. In this case each thread computed the primes only up to 10,000 (to make the runtime manageable). System time dominated, at 1546.968 seconds out of a total 2176.906 seconds (giving a user time of 629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite accurate. There is obviously no comparable test without the fix. Signed-off-by: Frank Mayhar <fmayhar@google.com> Cc: Roland McGrath <roland@redhat.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Cc: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-12 16:54:39 +00:00
static inline void
init_task_pid(struct task_struct *task, enum pid_type type, struct pid *pid)
{
task->pids[type].pid = pid;
}
static inline void rcu_copy_process(struct task_struct *p)
{
#ifdef CONFIG_PREEMPT_RCU
p->rcu_read_lock_nesting = 0;
p->rcu_read_unlock_special.s = 0;
p->rcu_blocked_node = NULL;
INIT_LIST_HEAD(&p->rcu_node_entry);
#endif /* #ifdef CONFIG_PREEMPT_RCU */
#ifdef CONFIG_TASKS_RCU
p->rcu_tasks_holdout = false;
INIT_LIST_HEAD(&p->rcu_tasks_holdout_list);
p->rcu_tasks_idle_cpu = -1;
#endif /* #ifdef CONFIG_TASKS_RCU */
}
/*
* This creates a new process as a copy of the old one,
* but does not actually start it yet.
*
* It copies the registers, and all the appropriate
* parts of the process environment (as per the clone
* flags). The actual kick-off is left to the caller.
*/
static __latent_entropy struct task_struct *copy_process(
unsigned long clone_flags,
unsigned long stack_start,
unsigned long stack_size,
int __user *child_tidptr,
struct pid *pid,
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
int trace,
unsigned long tls,
int node)
{
int retval;
struct task_struct *p;
if ((clone_flags & (CLONE_NEWNS|CLONE_FS)) == (CLONE_NEWNS|CLONE_FS))
return ERR_PTR(-EINVAL);
if ((clone_flags & (CLONE_NEWUSER|CLONE_FS)) == (CLONE_NEWUSER|CLONE_FS))
return ERR_PTR(-EINVAL);
/*
* Thread groups must share signals as well, and detached threads
* can only be started up within the thread group.
*/
if ((clone_flags & CLONE_THREAD) && !(clone_flags & CLONE_SIGHAND))
return ERR_PTR(-EINVAL);
/*
* Shared signal handlers imply shared VM. By way of the above,
* thread groups also imply shared VM. Blocking this case allows
* for various simplifications in other code.
*/
if ((clone_flags & CLONE_SIGHAND) && !(clone_flags & CLONE_VM))
return ERR_PTR(-EINVAL);
/*
* Siblings of global init remain as zombies on exit since they are
* not reaped by their parent (swapper). To solve this and to avoid
* multi-rooted process trees, prevent global and container-inits
* from creating siblings.
*/
if ((clone_flags & CLONE_PARENT) &&
current->signal->flags & SIGNAL_UNKILLABLE)
return ERR_PTR(-EINVAL);
/*
* If the new process will be in a different pid or user namespace
* do not allow it to share a thread group with the forking task.
*/
if (clone_flags & CLONE_THREAD) {
if ((clone_flags & (CLONE_NEWUSER | CLONE_NEWPID)) ||
(task_active_pid_ns(current) !=
current->nsproxy->pid_ns_for_children))
return ERR_PTR(-EINVAL);
}
retval = security_task_create(clone_flags);
if (retval)
goto fork_out;
retval = -ENOMEM;
p = dup_task_struct(current, node);
if (!p)
goto fork_out;
ftrace_graph_init_task(p);
rt_mutex_init_task(p);
#ifdef CONFIG_PROVE_LOCKING
DEBUG_LOCKS_WARN_ON(!p->hardirqs_enabled);
DEBUG_LOCKS_WARN_ON(!p->softirqs_enabled);
#endif
retval = -EAGAIN;
if (atomic_read(&p->real_cred->user->processes) >=
task_rlimit(p, RLIMIT_NPROC)) {
if (p->real_cred->user != INIT_USER &&
!capable(CAP_SYS_RESOURCE) && !capable(CAP_SYS_ADMIN))
goto bad_fork_free;
}
move RLIMIT_NPROC check from set_user() to do_execve_common() The patch http://lkml.org/lkml/2003/7/13/226 introduced an RLIMIT_NPROC check in set_user() to check for NPROC exceeding via setuid() and similar functions. Before the check there was a possibility to greatly exceed the allowed number of processes by an unprivileged user if the program relied on rlimit only. But the check created new security threat: many poorly written programs simply don't check setuid() return code and believe it cannot fail if executed with root privileges. So, the check is removed in this patch because of too often privilege escalations related to buggy programs. The NPROC can still be enforced in the common code flow of daemons spawning user processes. Most of daemons do fork()+setuid()+execve(). The check introduced in execve() (1) enforces the same limit as in setuid() and (2) doesn't create similar security issues. Neil Brown suggested to track what specific process has exceeded the limit by setting PF_NPROC_EXCEEDED process flag. With the change only this process would fail on execve(), and other processes' execve() behaviour is not changed. Solar Designer suggested to re-check whether NPROC limit is still exceeded at the moment of execve(). If the process was sleeping for days between set*uid() and execve(), and the NPROC counter step down under the limit, the defered execve() failure because NPROC limit was exceeded days ago would be unexpected. If the limit is not exceeded anymore, we clear the flag on successful calls to execve() and fork(). The flag is also cleared on successful calls to set_user() as the limit was exceeded for the previous user, not the current one. Similar check was introduced in -ow patches (without the process flag). v3 - clear PF_NPROC_EXCEEDED on successful calls to set_user(). Reviewed-by: James Morris <jmorris@namei.org> Signed-off-by: Vasiliy Kulikov <segoon@openwall.com> Acked-by: NeilBrown <neilb@suse.de> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-08-08 15:02:04 +00:00
current->flags &= ~PF_NPROC_EXCEEDED;
retval = copy_creds(p, clone_flags);
if (retval < 0)
goto bad_fork_free;
/*
* If multiple threads are within copy_process(), then this check
* triggers too late. This doesn't hurt, the check is only there
* to stop root fork bombs.
*/
retval = -EAGAIN;
if (nr_threads >= max_threads)
goto bad_fork_cleanup_count;
delayacct_tsk_init(p); /* Must remain after dup_task_struct() */
p->flags &= ~(PF_SUPERPRIV | PF_WQ_WORKER | PF_IDLE);
p->flags |= PF_FORKNOEXEC;
INIT_LIST_HEAD(&p->children);
INIT_LIST_HEAD(&p->sibling);
rcu: Merge preemptable-RCU functionality into hierarchical RCU Create a kernel/rcutree_plugin.h file that contains definitions for preemptable RCU (or, under the #else branch of the #ifdef, empty definitions for the classic non-preemptable semantics). These definitions fit into plugins defined in kernel/rcutree.c for this purpose. This variant of preemptable RCU uses a new algorithm whose read-side expense is roughly that of classic hierarchical RCU under CONFIG_PREEMPT. This new algorithm's update-side expense is similar to that of classic hierarchical RCU, and, in absence of read-side preemption or blocking, is exactly that of classic hierarchical RCU. Perhaps more important, this new algorithm has a much simpler implementation, saving well over 1,000 lines of code compared to mainline's implementation of preemptable RCU, which will hopefully be retired in favor of this new algorithm. The simplifications are obtained by maintaining per-task nesting state for running tasks, and using a simple lock-protected algorithm to handle accounting when tasks block within RCU read-side critical sections, making use of lessons learned while creating numerous user-level RCU implementations over the past 18 months. Signed-off-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com> Cc: laijs@cn.fujitsu.com Cc: dipankar@in.ibm.com Cc: akpm@linux-foundation.org Cc: mathieu.desnoyers@polymtl.ca Cc: josht@linux.vnet.ibm.com Cc: dvhltc@us.ibm.com Cc: niv@us.ibm.com Cc: peterz@infradead.org Cc: rostedt@goodmis.org LKML-Reference: <12509746134003-git-send-email-> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-08-22 20:56:52 +00:00
rcu_copy_process(p);
p->vfork_done = NULL;
spin_lock_init(&p->alloc_lock);
init_sigpending(&p->pending);
p->utime = p->stime = p->gtime = 0;
#ifdef CONFIG_ARCH_HAS_SCALED_CPUTIME
p->utimescaled = p->stimescaled = 0;
#endif
prev_cputime_init(&p->prev_cputime);
#ifdef CONFIG_VIRT_CPU_ACCOUNTING_GEN
seqcount_init(&p->vtime_seqcount);
p->vtime_snap = 0;
p->vtime_snap_whence = VTIME_INACTIVE;
#endif
#if defined(SPLIT_RSS_COUNTING)
memset(&p->rss_stat, 0, sizeof(p->rss_stat));
#endif
p->default_timer_slack_ns = current->timer_slack_ns;
task_io_accounting_init(&p->ioac);
acct_clear_integrals(p);
timers: fix itimer/many thread hang Overview This patch reworks the handling of POSIX CPU timers, including the ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together with the help of Roland McGrath, the owner and original writer of this code. The problem we ran into, and the reason for this rework, has to do with using a profiling timer in a process with a large number of threads. It appears that the performance of the old implementation of run_posix_cpu_timers() was at least O(n*3) (where "n" is the number of threads in a process) or worse. Everything is fine with an increasing number of threads until the time taken for that routine to run becomes the same as or greater than the tick time, at which point things degrade rather quickly. This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF." Code Changes This rework corrects the implementation of run_posix_cpu_timers() to make it run in constant time for a particular machine. (Performance may vary between one machine and another depending upon whether the kernel is built as single- or multiprocessor and, in the latter case, depending upon the number of running processors.) To do this, at each tick we now update fields in signal_struct as well as task_struct. The run_posix_cpu_timers() function uses those fields to make its decisions. We define a new structure, "task_cputime," to contain user, system and scheduler times and use these in appropriate places: struct task_cputime { cputime_t utime; cputime_t stime; unsigned long long sum_exec_runtime; }; This is included in the structure "thread_group_cputime," which is a new substructure of signal_struct and which varies for uniprocessor versus multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as a simple substructure, while for multiprocessor kernels it is a pointer: struct thread_group_cputime { struct task_cputime totals; }; struct thread_group_cputime { struct task_cputime *totals; }; We also add a new task_cputime substructure directly to signal_struct, to cache the earliest expiration of process-wide timers, and task_cputime also replaces the it_*_expires fields of task_struct (used for earliest expiration of thread timers). The "thread_group_cputime" structure contains process-wide timers that are updated via account_user_time() and friends. In the non-SMP case the structure is a simple aggregator; unfortunately in the SMP case that simplicity was not achievable due to cache-line contention between CPUs (in one measured case performance was actually _worse_ on a 16-cpu system than the same test on a 4-cpu system, due to this contention). For SMP, the thread_group_cputime counters are maintained as a per-cpu structure allocated using alloc_percpu(). The timer functions update only the timer field in the structure corresponding to the running CPU, obtained using per_cpu_ptr(). We define a set of inline functions in sched.h that we use to maintain the thread_group_cputime structure and hide the differences between UP and SMP implementations from the rest of the kernel. The thread_group_cputime_init() function initializes the thread_group_cputime structure for the given task. The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the out-of-line function thread_group_cputime_alloc_smp() to allocate and fill in the per-cpu structures and fields. The thread_group_cputime_free() function, also a no-op for UP, in SMP frees the per-cpu structures. The thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls thread_group_cputime_alloc() if the per-cpu structures haven't yet been allocated. The thread_group_cputime() function fills the task_cputime structure it is passed with the contents of the thread_group_cputime fields; in UP it's that simple but in SMP it must also safely check that tsk->signal is non-NULL (if it is it just uses the appropriate fields of task_struct) and, if so, sums the per-cpu values for each online CPU. Finally, the three functions account_group_user_time(), account_group_system_time() and account_group_exec_runtime() are used by timer functions to update the respective fields of the thread_group_cputime structure. Non-SMP operation is trivial and will not be mentioned further. The per-cpu structure is always allocated when a task creates its first new thread, via a call to thread_group_cputime_clone_thread() from copy_signal(). It is freed at process exit via a call to thread_group_cputime_free() from cleanup_signal(). All functions that formerly summed utime/stime/sum_sched_runtime values from from all threads in the thread group now use thread_group_cputime() to snapshot the values in the thread_group_cputime structure or the values in the task structure itself if the per-cpu structure hasn't been allocated. Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit. The run_posix_cpu_timers() function has been split into a fast path and a slow path; the former safely checks whether there are any expired thread timers and, if not, just returns, while the slow path does the heavy lifting. With the dedicated thread group fields, timers are no longer "rebalanced" and the process_timer_rebalance() function and related code has gone away. All summing loops are gone and all code that used them now uses the thread_group_cputime() inline. When process-wide timers are set, the new task_cputime structure in signal_struct is used to cache the earliest expiration; this is checked in the fast path. Performance The fix appears not to add significant overhead to existing operations. It generally performs the same as the current code except in two cases, one in which it performs slightly worse (Case 5 below) and one in which it performs very significantly better (Case 2 below). Overall it's a wash except in those two cases. I've since done somewhat more involved testing on a dual-core Opteron system. Case 1: With no itimer running, for a test with 100,000 threads, the fixed kernel took 1428.5 seconds, 513 seconds more than the unfixed system, all of which was spent in the system. There were twice as many voluntary context switches with the fix as without it. Case 2: With an itimer running at .01 second ticks and 4000 threads (the most an unmodified kernel can handle), the fixed kernel ran the test in eight percent of the time (5.8 seconds as opposed to 70 seconds) and had better tick accuracy (.012 seconds per tick as opposed to .023 seconds per tick). Case 3: A 4000-thread test with an initial timer tick of .01 second and an interval of 10,000 seconds (i.e. a timer that ticks only once) had very nearly the same performance in both cases: 6.3 seconds elapsed for the fixed kernel versus 5.5 seconds for the unfixed kernel. With fewer threads (eight in these tests), the Case 1 test ran in essentially the same time on both the modified and unmodified kernels (5.2 seconds versus 5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds versus 5.4 seconds but again with much better tick accuracy, .013 seconds per tick versus .025 seconds per tick for the unmodified kernel. Since the fix affected the rlimit code, I also tested soft and hard CPU limits. Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer running), the modified kernel was very slightly favored in that while it killed the process in 19.997 seconds of CPU time (5.002 seconds of wall time), only .003 seconds of that was system time, the rest was user time. The unmodified kernel killed the process in 20.001 seconds of CPU (5.014 seconds of wall time) of which .016 seconds was system time. Really, though, the results were too close to call. The results were essentially the same with no itimer running. Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds (where the hard limit would never be reached) and an itimer running, the modified kernel exhibited worse tick accuracy than the unmodified kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise, performance was almost indistinguishable. With no itimer running this test exhibited virtually identical behavior and times in both cases. In times past I did some limited performance testing. those results are below. On a four-cpu Opteron system without this fix, a sixteen-thread test executed in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On the same system with the fix, user and elapsed time were about the same, but system time dropped to 0.007 seconds. Performance with eight, four and one thread were comparable. Interestingly, the timer ticks with the fix seemed more accurate: The sixteen-thread test with the fix received 149543 ticks for 0.024 seconds per tick, while the same test without the fix received 58720 for 0.061 seconds per tick. Both cases were configured for an interval of 0.01 seconds. Again, the other tests were comparable. Each thread in this test computed the primes up to 25,000,000. I also did a test with a large number of threads, 100,000 threads, which is impossible without the fix. In this case each thread computed the primes only up to 10,000 (to make the runtime manageable). System time dominated, at 1546.968 seconds out of a total 2176.906 seconds (giving a user time of 629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite accurate. There is obviously no comparable test without the fix. Signed-off-by: Frank Mayhar <fmayhar@google.com> Cc: Roland McGrath <roland@redhat.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Cc: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-12 16:54:39 +00:00
posix_cpu_timers_init(p);
p->start_time = ktime_get_ns();
p->real_start_time = ktime_get_boot_ns();
p->io_context = NULL;
p->audit_context = NULL;
cgroup_fork(p);
#ifdef CONFIG_NUMA
p->mempolicy = mpol_dup(p->mempolicy);
if (IS_ERR(p->mempolicy)) {
retval = PTR_ERR(p->mempolicy);
p->mempolicy = NULL;
goto bad_fork_cleanup_threadgroup_lock;
}
#endif
cpusets: randomize node rotor used in cpuset_mem_spread_node() [ This patch has already been accepted as commit 0ac0c0d0f837 but later reverted (commit 35926ff5fba8) because it itroduced arch specific __node_random which was defined only for x86 code so it broke other archs. This is a followup without any arch specific code. Other than that there are no functional changes.] Some workloads that create a large number of small files tend to assign too many pages to node 0 (multi-node systems). Part of the reason is that the rotor (in cpuset_mem_spread_node()) used to assign nodes starts at node 0 for newly created tasks. This patch changes the rotor to be initialized to a random node number of the cpuset. [akpm@linux-foundation.org: fix layout] [Lee.Schermerhorn@hp.com: Define stub numa_random() for !NUMA configuration] [mhocko@suse.cz: Make it arch independent] [akpm@linux-foundation.org: fix CONFIG_NUMA=y, MAX_NUMNODES>1 build] Signed-off-by: Jack Steiner <steiner@sgi.com> Signed-off-by: Lee Schermerhorn <lee.schermerhorn@hp.com> Signed-off-by: Michal Hocko <mhocko@suse.cz> Reviewed-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Christoph Lameter <cl@linux-foundation.org> Cc: Pekka Enberg <penberg@cs.helsinki.fi> Cc: Paul Menage <menage@google.com> Cc: Jack Steiner <steiner@sgi.com> Cc: Robin Holt <holt@sgi.com> Cc: David Rientjes <rientjes@google.com> Cc: Christoph Lameter <cl@linux-foundation.org> Cc: David Rientjes <rientjes@google.com> Cc: Jack Steiner <steiner@sgi.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Lee Schermerhorn <lee.schermerhorn@hp.com> Cc: Michal Hocko <mhocko@suse.cz> Cc: Paul Menage <menage@google.com> Cc: Pekka Enberg <penberg@cs.helsinki.fi> Cc: Robin Holt <holt@sgi.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-07-26 23:08:30 +00:00
#ifdef CONFIG_CPUSETS
p->cpuset_mem_spread_rotor = NUMA_NO_NODE;
p->cpuset_slab_spread_rotor = NUMA_NO_NODE;
cpuset: mm: reduce large amounts of memory barrier related damage v3 Commit c0ff7453bb5c ("cpuset,mm: fix no node to alloc memory when changing cpuset's mems") wins a super prize for the largest number of memory barriers entered into fast paths for one commit. [get|put]_mems_allowed is incredibly heavy with pairs of full memory barriers inserted into a number of hot paths. This was detected while investigating at large page allocator slowdown introduced some time after 2.6.32. The largest portion of this overhead was shown by oprofile to be at an mfence introduced by this commit into the page allocator hot path. For extra style points, the commit introduced the use of yield() in an implementation of what looks like a spinning mutex. This patch replaces the full memory barriers on both read and write sides with a sequence counter with just read barriers on the fast path side. This is much cheaper on some architectures, including x86. The main bulk of the patch is the retry logic if the nodemask changes in a manner that can cause a false failure. While updating the nodemask, a check is made to see if a false failure is a risk. If it is, the sequence number gets bumped and parallel allocators will briefly stall while the nodemask update takes place. In a page fault test microbenchmark, oprofile samples from __alloc_pages_nodemask went from 4.53% of all samples to 1.15%. The actual results were 3.3.0-rc3 3.3.0-rc3 rc3-vanilla nobarrier-v2r1 Clients 1 UserTime 0.07 ( 0.00%) 0.08 (-14.19%) Clients 2 UserTime 0.07 ( 0.00%) 0.07 ( 2.72%) Clients 4 UserTime 0.08 ( 0.00%) 0.07 ( 3.29%) Clients 1 SysTime 0.70 ( 0.00%) 0.65 ( 6.65%) Clients 2 SysTime 0.85 ( 0.00%) 0.82 ( 3.65%) Clients 4 SysTime 1.41 ( 0.00%) 1.41 ( 0.32%) Clients 1 WallTime 0.77 ( 0.00%) 0.74 ( 4.19%) Clients 2 WallTime 0.47 ( 0.00%) 0.45 ( 3.73%) Clients 4 WallTime 0.38 ( 0.00%) 0.37 ( 1.58%) Clients 1 Flt/sec/cpu 497620.28 ( 0.00%) 520294.53 ( 4.56%) Clients 2 Flt/sec/cpu 414639.05 ( 0.00%) 429882.01 ( 3.68%) Clients 4 Flt/sec/cpu 257959.16 ( 0.00%) 258761.48 ( 0.31%) Clients 1 Flt/sec 495161.39 ( 0.00%) 517292.87 ( 4.47%) Clients 2 Flt/sec 820325.95 ( 0.00%) 850289.77 ( 3.65%) Clients 4 Flt/sec 1020068.93 ( 0.00%) 1022674.06 ( 0.26%) MMTests Statistics: duration Sys Time Running Test (seconds) 135.68 132.17 User+Sys Time Running Test (seconds) 164.2 160.13 Total Elapsed Time (seconds) 123.46 120.87 The overall improvement is small but the System CPU time is much improved and roughly in correlation to what oprofile reported (these performance figures are without profiling so skew is expected). The actual number of page faults is noticeably improved. For benchmarks like kernel builds, the overall benefit is marginal but the system CPU time is slightly reduced. To test the actual bug the commit fixed I opened two terminals. The first ran within a cpuset and continually ran a small program that faulted 100M of anonymous data. In a second window, the nodemask of the cpuset was continually randomised in a loop. Without the commit, the program would fail every so often (usually within 10 seconds) and obviously with the commit everything worked fine. With this patch applied, it also worked fine so the fix should be functionally equivalent. Signed-off-by: Mel Gorman <mgorman@suse.de> Cc: Miao Xie <miaox@cn.fujitsu.com> Cc: David Rientjes <rientjes@google.com> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: Christoph Lameter <cl@linux.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-03-21 23:34:11 +00:00
seqcount_init(&p->mems_allowed_seq);
cpusets: randomize node rotor used in cpuset_mem_spread_node() [ This patch has already been accepted as commit 0ac0c0d0f837 but later reverted (commit 35926ff5fba8) because it itroduced arch specific __node_random which was defined only for x86 code so it broke other archs. This is a followup without any arch specific code. Other than that there are no functional changes.] Some workloads that create a large number of small files tend to assign too many pages to node 0 (multi-node systems). Part of the reason is that the rotor (in cpuset_mem_spread_node()) used to assign nodes starts at node 0 for newly created tasks. This patch changes the rotor to be initialized to a random node number of the cpuset. [akpm@linux-foundation.org: fix layout] [Lee.Schermerhorn@hp.com: Define stub numa_random() for !NUMA configuration] [mhocko@suse.cz: Make it arch independent] [akpm@linux-foundation.org: fix CONFIG_NUMA=y, MAX_NUMNODES>1 build] Signed-off-by: Jack Steiner <steiner@sgi.com> Signed-off-by: Lee Schermerhorn <lee.schermerhorn@hp.com> Signed-off-by: Michal Hocko <mhocko@suse.cz> Reviewed-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Christoph Lameter <cl@linux-foundation.org> Cc: Pekka Enberg <penberg@cs.helsinki.fi> Cc: Paul Menage <menage@google.com> Cc: Jack Steiner <steiner@sgi.com> Cc: Robin Holt <holt@sgi.com> Cc: David Rientjes <rientjes@google.com> Cc: Christoph Lameter <cl@linux-foundation.org> Cc: David Rientjes <rientjes@google.com> Cc: Jack Steiner <steiner@sgi.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Lee Schermerhorn <lee.schermerhorn@hp.com> Cc: Michal Hocko <mhocko@suse.cz> Cc: Paul Menage <menage@google.com> Cc: Pekka Enberg <penberg@cs.helsinki.fi> Cc: Robin Holt <holt@sgi.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-07-26 23:08:30 +00:00
#endif
#ifdef CONFIG_TRACE_IRQFLAGS
p->irq_events = 0;
p->hardirqs_enabled = 0;
p->hardirq_enable_ip = 0;
p->hardirq_enable_event = 0;
p->hardirq_disable_ip = _THIS_IP_;
p->hardirq_disable_event = 0;
p->softirqs_enabled = 1;
p->softirq_enable_ip = _THIS_IP_;
p->softirq_enable_event = 0;
p->softirq_disable_ip = 0;
p->softirq_disable_event = 0;
p->hardirq_context = 0;
p->softirq_context = 0;
#endif
sched/preempt, mm/fault: Count pagefault_disable() levels in pagefault_disabled Until now, pagefault_disable()/pagefault_enabled() used the preempt count to track whether in an environment with pagefaults disabled (can be queried via in_atomic()). This patch introduces a separate counter in task_struct to count the level of pagefault_disable() calls. We'll keep manipulating the preempt count to retain compatibility to existing pagefault handlers. It is now possible to verify whether in a pagefault_disable() envionment by calling pagefault_disabled(). In contrast to in_atomic() it will not be influenced by preempt_enable()/preempt_disable(). This patch is based on a patch from Ingo Molnar. Reviewed-and-tested-by: Thomas Gleixner <tglx@linutronix.de> Signed-off-by: David Hildenbrand <dahi@linux.vnet.ibm.com> Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org> Cc: David.Laight@ACULAB.COM Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: airlied@linux.ie Cc: akpm@linux-foundation.org Cc: benh@kernel.crashing.org Cc: bigeasy@linutronix.de Cc: borntraeger@de.ibm.com Cc: daniel.vetter@intel.com Cc: heiko.carstens@de.ibm.com Cc: herbert@gondor.apana.org.au Cc: hocko@suse.cz Cc: hughd@google.com Cc: mst@redhat.com Cc: paulus@samba.org Cc: ralf@linux-mips.org Cc: schwidefsky@de.ibm.com Cc: yang.shi@windriver.com Link: http://lkml.kernel.org/r/1431359540-32227-2-git-send-email-dahi@linux.vnet.ibm.com Signed-off-by: Ingo Molnar <mingo@kernel.org>
2015-05-11 15:52:06 +00:00
p->pagefault_disabled = 0;
[PATCH] lockdep: core Do 'make oldconfig' and accept all the defaults for new config options - reboot into the kernel and if everything goes well it should boot up fine and you should have /proc/lockdep and /proc/lockdep_stats files. Typically if the lock validator finds some problem it will print out voluminous debug output that begins with "BUG: ..." and which syslog output can be used by kernel developers to figure out the precise locking scenario. What does the lock validator do? It "observes" and maps all locking rules as they occur dynamically (as triggered by the kernel's natural use of spinlocks, rwlocks, mutexes and rwsems). Whenever the lock validator subsystem detects a new locking scenario, it validates this new rule against the existing set of rules. If this new rule is consistent with the existing set of rules then the new rule is added transparently and the kernel continues as normal. If the new rule could create a deadlock scenario then this condition is printed out. When determining validity of locking, all possible "deadlock scenarios" are considered: assuming arbitrary number of CPUs, arbitrary irq context and task context constellations, running arbitrary combinations of all the existing locking scenarios. In a typical system this means millions of separate scenarios. This is why we call it a "locking correctness" validator - for all rules that are observed the lock validator proves it with mathematical certainty that a deadlock could not occur (assuming that the lock validator implementation itself is correct and its internal data structures are not corrupted by some other kernel subsystem). [see more details and conditionals of this statement in include/linux/lockdep.h and Documentation/lockdep-design.txt] Furthermore, this "all possible scenarios" property of the validator also enables the finding of complex, highly unlikely multi-CPU multi-context races via single single-context rules, increasing the likelyhood of finding bugs drastically. In practical terms: the lock validator already found a bug in the upstream kernel that could only occur on systems with 3 or more CPUs, and which needed 3 very unlikely code sequences to occur at once on the 3 CPUs. That bug was found and reported on a single-CPU system (!). So in essence a race will be found "piecemail-wise", triggering all the necessary components for the race, without having to reproduce the race scenario itself! In its short existence the lock validator found and reported many bugs before they actually caused a real deadlock. To further increase the efficiency of the validator, the mapping is not per "lock instance", but per "lock-class". For example, all struct inode objects in the kernel have inode->inotify_mutex. If there are 10,000 inodes cached, then there are 10,000 lock objects. But ->inotify_mutex is a single "lock type", and all locking activities that occur against ->inotify_mutex are "unified" into this single lock-class. The advantage of the lock-class approach is that all historical ->inotify_mutex uses are mapped into a single (and as narrow as possible) set of locking rules - regardless of how many different tasks or inode structures it took to build this set of rules. The set of rules persist during the lifetime of the kernel. To see the rough magnitude of checking that the lock validator does, here's a portion of /proc/lockdep_stats, fresh after bootup: lock-classes: 694 [max: 2048] direct dependencies: 1598 [max: 8192] indirect dependencies: 17896 all direct dependencies: 16206 dependency chains: 1910 [max: 8192] in-hardirq chains: 17 in-softirq chains: 105 in-process chains: 1065 stack-trace entries: 38761 [max: 131072] combined max dependencies: 2033928 hardirq-safe locks: 24 hardirq-unsafe locks: 176 softirq-safe locks: 53 softirq-unsafe locks: 137 irq-safe locks: 59 irq-unsafe locks: 176 The lock validator has observed 1598 actual single-thread locking patterns, and has validated all possible 2033928 distinct locking scenarios. More details about the design of the lock validator can be found in Documentation/lockdep-design.txt, which can also found at: http://redhat.com/~mingo/lockdep-patches/lockdep-design.txt [bunk@stusta.de: cleanups] Signed-off-by: Ingo Molnar <mingo@elte.hu> Signed-off-by: Arjan van de Ven <arjan@linux.intel.com> Signed-off-by: Adrian Bunk <bunk@stusta.de> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-07-03 07:24:50 +00:00
#ifdef CONFIG_LOCKDEP
p->lockdep_depth = 0; /* no locks held yet */
p->curr_chain_key = 0;
p->lockdep_recursion = 0;
#endif
#ifdef CONFIG_DEBUG_MUTEXES
p->blocked_on = NULL; /* not blocked yet */
#endif
#ifdef CONFIG_BCACHE
p->sequential_io = 0;
p->sequential_io_avg = 0;
#endif
sched: fix copy_namespace() <-> sched_fork() dependency in do_fork Sukadev Bhattiprolu reported a kernel crash with control groups. There are couple of problems discovered by Suka's test: - The test requires the cgroup filesystem to be mounted with atleast the cpu and ns options (i.e both namespace and cpu controllers are active in the same hierarchy). # mkdir /dev/cpuctl # mount -t cgroup -ocpu,ns none cpuctl (or simply) # mount -t cgroup none cpuctl -> Will activate all controllers in same hierarchy. - The test invokes clone() with CLONE_NEWNS set. This causes a a new child to be created, also a new group (do_fork->copy_namespaces->ns_cgroup_clone-> cgroup_clone) and the child is attached to the new group (cgroup_clone-> attach_task->sched_move_task). At this point in time, the child's scheduler related fields are uninitialized (including its on_rq field, which it has inherited from parent). As a result sched_move_task thinks its on runqueue, when it isn't. As a solution to this problem, I moved sched_fork() call, which initializes scheduler related fields on a new task, before copy_namespaces(). I am not sure though whether moving up will cause other side-effects. Do you see any issue? - The second problem exposed by this test is that task_new_fair() assumes that parent and child will be part of the same group (which needn't be as this test shows). As a result, cfs_rq->curr can be NULL for the child. The solution is to test for curr pointer being NULL in task_new_fair(). With the patch below, I could run ns_exec() fine w/o a crash. Reported-by: Sukadev Bhattiprolu <sukadev@us.ibm.com> Signed-off-by: Srivatsa Vaddagiri <vatsa@linux.vnet.ibm.com> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2007-11-09 21:39:39 +00:00
/* Perform scheduler related setup. Assign this task to a CPU. */
sched/deadline: Add SCHED_DEADLINE structures & implementation Introduces the data structures, constants and symbols needed for SCHED_DEADLINE implementation. Core data structure of SCHED_DEADLINE are defined, along with their initializers. Hooks for checking if a task belong to the new policy are also added where they are needed. Adds a scheduling class, in sched/dl.c and a new policy called SCHED_DEADLINE. It is an implementation of the Earliest Deadline First (EDF) scheduling algorithm, augmented with a mechanism (called Constant Bandwidth Server, CBS) that makes it possible to isolate the behaviour of tasks between each other. The typical -deadline task will be made up of a computation phase (instance) which is activated on a periodic or sporadic fashion. The expected (maximum) duration of such computation is called the task's runtime; the time interval by which each instance need to be completed is called the task's relative deadline. The task's absolute deadline is dynamically calculated as the time instant a task (better, an instance) activates plus the relative deadline. The EDF algorithms selects the task with the smallest absolute deadline as the one to be executed first, while the CBS ensures each task to run for at most its runtime every (relative) deadline length time interval, avoiding any interference between different tasks (bandwidth isolation). Thanks to this feature, also tasks that do not strictly comply with the computational model sketched above can effectively use the new policy. To summarize, this patch: - introduces the data structures, constants and symbols needed; - implements the core logic of the scheduling algorithm in the new scheduling class file; - provides all the glue code between the new scheduling class and the core scheduler and refines the interactions between sched/dl and the other existing scheduling classes. Signed-off-by: Dario Faggioli <raistlin@linux.it> Signed-off-by: Michael Trimarchi <michael@amarulasolutions.com> Signed-off-by: Fabio Checconi <fchecconi@gmail.com> Signed-off-by: Juri Lelli <juri.lelli@gmail.com> Signed-off-by: Peter Zijlstra <peterz@infradead.org> Link: http://lkml.kernel.org/r/1383831828-15501-4-git-send-email-juri.lelli@gmail.com Signed-off-by: Ingo Molnar <mingo@kernel.org>
2013-11-28 10:14:43 +00:00
retval = sched_fork(clone_flags, p);
if (retval)
goto bad_fork_cleanup_policy;
perf: Do the big rename: Performance Counters -> Performance Events Bye-bye Performance Counters, welcome Performance Events! In the past few months the perfcounters subsystem has grown out its initial role of counting hardware events, and has become (and is becoming) a much broader generic event enumeration, reporting, logging, monitoring, analysis facility. Naming its core object 'perf_counter' and naming the subsystem 'perfcounters' has become more and more of a misnomer. With pending code like hw-breakpoints support the 'counter' name is less and less appropriate. All in one, we've decided to rename the subsystem to 'performance events' and to propagate this rename through all fields, variables and API names. (in an ABI compatible fashion) The word 'event' is also a bit shorter than 'counter' - which makes it slightly more convenient to write/handle as well. Thanks goes to Stephane Eranian who first observed this misnomer and suggested a rename. User-space tooling and ABI compatibility is not affected - this patch should be function-invariant. (Also, defconfigs were not touched to keep the size down.) This patch has been generated via the following script: FILES=$(find * -type f | grep -vE 'oprofile|[^K]config') sed -i \ -e 's/PERF_EVENT_/PERF_RECORD_/g' \ -e 's/PERF_COUNTER/PERF_EVENT/g' \ -e 's/perf_counter/perf_event/g' \ -e 's/nb_counters/nb_events/g' \ -e 's/swcounter/swevent/g' \ -e 's/tpcounter_event/tp_event/g' \ $FILES for N in $(find . -name perf_counter.[ch]); do M=$(echo $N | sed 's/perf_counter/perf_event/g') mv $N $M done FILES=$(find . -name perf_event.*) sed -i \ -e 's/COUNTER_MASK/REG_MASK/g' \ -e 's/COUNTER/EVENT/g' \ -e 's/\<event\>/event_id/g' \ -e 's/counter/event/g' \ -e 's/Counter/Event/g' \ $FILES ... to keep it as correct as possible. This script can also be used by anyone who has pending perfcounters patches - it converts a Linux kernel tree over to the new naming. We tried to time this change to the point in time where the amount of pending patches is the smallest: the end of the merge window. Namespace clashes were fixed up in a preparatory patch - and some stylistic fallout will be fixed up in a subsequent patch. ( NOTE: 'counters' are still the proper terminology when we deal with hardware registers - and these sed scripts are a bit over-eager in renaming them. I've undone some of that, but in case there's something left where 'counter' would be better than 'event' we can undo that on an individual basis instead of touching an otherwise nicely automated patch. ) Suggested-by: Stephane Eranian <eranian@google.com> Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl> Acked-by: Paul Mackerras <paulus@samba.org> Reviewed-by: Arjan van de Ven <arjan@linux.intel.com> Cc: Mike Galbraith <efault@gmx.de> Cc: Arnaldo Carvalho de Melo <acme@redhat.com> Cc: Frederic Weisbecker <fweisbec@gmail.com> Cc: Steven Rostedt <rostedt@goodmis.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: David Howells <dhowells@redhat.com> Cc: Kyle McMartin <kyle@mcmartin.ca> Cc: Martin Schwidefsky <schwidefsky@de.ibm.com> Cc: "David S. Miller" <davem@davemloft.net> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: <linux-arch@vger.kernel.org> LKML-Reference: <new-submission> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-09-21 10:02:48 +00:00
retval = perf_event_init_task(p);
if (retval)
goto bad_fork_cleanup_policy;
retval = audit_alloc(p);
if (retval)
goto bad_fork_cleanup_perf;
/* copy all the process information */
shm: make exit_shm work proportional to task activity This is small set of patches our team has had kicking around for a few versions internally that fixes tasks getting hung on shm_exit when there are many threads hammering it at once. Anton wrote a simple test to cause the issue: http://ozlabs.org/~anton/junkcode/bust_shm_exit.c Before applying this patchset, this test code will cause either hanging tracebacks or pthread out of memory errors. After this patchset, it will still produce output like: root@somehost:~# ./bust_shm_exit 1024 160 ... INFO: rcu_sched detected stalls on CPUs/tasks: {} (detected by 116, t=2111 jiffies, g=241, c=240, q=7113) INFO: Stall ended before state dump start ... But the task will continue to run along happily, so we consider this an improvement over hanging, even if it's a bit noisy. This patch (of 3): exit_shm obtains the ipc_ns shm rwsem for write and holds it while it walks every shared memory segment in the namespace. Thus the amount of work is related to the number of shm segments in the namespace not the number of segments that might need to be cleaned. In addition, this occurs after the task has been notified the thread has exited, so the number of tasks waiting for the ns shm rwsem can grow without bound until memory is exausted. Add a list to the task struct of all shmids allocated by this task. Init the list head in copy_process. Use the ns->rwsem for locking. Add segments after id is added, remove before removing from id. On unshare of NEW_IPCNS orphan any ids as if the task had exited, similar to handling of semaphore undo. I chose a define for the init sequence since its a simple list init, otherwise it would require a function call to avoid include loops between the semaphore code and the task struct. Converting the list_del to list_del_init for the unshare cases would remove the exit followed by init, but I left it blow up if not inited. Signed-off-by: Milton Miller <miltonm@bga.com> Signed-off-by: Jack Miller <millerjo@us.ibm.com> Cc: Davidlohr Bueso <davidlohr@hp.com> Cc: Manfred Spraul <manfred@colorfullife.com> Cc: Anton Blanchard <anton@samba.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-08 21:23:19 +00:00
shm_init_task(p);
retval = copy_semundo(clone_flags, p);
if (retval)
goto bad_fork_cleanup_audit;
retval = copy_files(clone_flags, p);
if (retval)
goto bad_fork_cleanup_semundo;
retval = copy_fs(clone_flags, p);
if (retval)
goto bad_fork_cleanup_files;
retval = copy_sighand(clone_flags, p);
if (retval)
goto bad_fork_cleanup_fs;
retval = copy_signal(clone_flags, p);
if (retval)
goto bad_fork_cleanup_sighand;
retval = copy_mm(clone_flags, p);
if (retval)
goto bad_fork_cleanup_signal;
retval = copy_namespaces(clone_flags, p);
if (retval)
CRED: Inaugurate COW credentials Inaugurate copy-on-write credentials management. This uses RCU to manage the credentials pointer in the task_struct with respect to accesses by other tasks. A process may only modify its own credentials, and so does not need locking to access or modify its own credentials. A mutex (cred_replace_mutex) is added to the task_struct to control the effect of PTRACE_ATTACHED on credential calculations, particularly with respect to execve(). With this patch, the contents of an active credentials struct may not be changed directly; rather a new set of credentials must be prepared, modified and committed using something like the following sequence of events: struct cred *new = prepare_creds(); int ret = blah(new); if (ret < 0) { abort_creds(new); return ret; } return commit_creds(new); There are some exceptions to this rule: the keyrings pointed to by the active credentials may be instantiated - keyrings violate the COW rule as managing COW keyrings is tricky, given that it is possible for a task to directly alter the keys in a keyring in use by another task. To help enforce this, various pointers to sets of credentials, such as those in the task_struct, are declared const. The purpose of this is compile-time discouragement of altering credentials through those pointers. Once a set of credentials has been made public through one of these pointers, it may not be modified, except under special circumstances: (1) Its reference count may incremented and decremented. (2) The keyrings to which it points may be modified, but not replaced. The only safe way to modify anything else is to create a replacement and commit using the functions described in Documentation/credentials.txt (which will be added by a later patch). This patch and the preceding patches have been tested with the LTP SELinux testsuite. This patch makes several logical sets of alteration: (1) execve(). This now prepares and commits credentials in various places in the security code rather than altering the current creds directly. (2) Temporary credential overrides. do_coredump() and sys_faccessat() now prepare their own credentials and temporarily override the ones currently on the acting thread, whilst preventing interference from other threads by holding cred_replace_mutex on the thread being dumped. This will be replaced in a future patch by something that hands down the credentials directly to the functions being called, rather than altering the task's objective credentials. (3) LSM interface. A number of functions have been changed, added or removed: (*) security_capset_check(), ->capset_check() (*) security_capset_set(), ->capset_set() Removed in favour of security_capset(). (*) security_capset(), ->capset() New. This is passed a pointer to the new creds, a pointer to the old creds and the proposed capability sets. It should fill in the new creds or return an error. All pointers, barring the pointer to the new creds, are now const. (*) security_bprm_apply_creds(), ->bprm_apply_creds() Changed; now returns a value, which will cause the process to be killed if it's an error. (*) security_task_alloc(), ->task_alloc_security() Removed in favour of security_prepare_creds(). (*) security_cred_free(), ->cred_free() New. Free security data attached to cred->security. (*) security_prepare_creds(), ->cred_prepare() New. Duplicate any security data attached to cred->security. (*) security_commit_creds(), ->cred_commit() New. Apply any security effects for the upcoming installation of new security by commit_creds(). (*) security_task_post_setuid(), ->task_post_setuid() Removed in favour of security_task_fix_setuid(). (*) security_task_fix_setuid(), ->task_fix_setuid() Fix up the proposed new credentials for setuid(). This is used by cap_set_fix_setuid() to implicitly adjust capabilities in line with setuid() changes. Changes are made to the new credentials, rather than the task itself as in security_task_post_setuid(). (*) security_task_reparent_to_init(), ->task_reparent_to_init() Removed. Instead the task being reparented to init is referred directly to init's credentials. NOTE! This results in the loss of some state: SELinux's osid no longer records the sid of the thread that forked it. (*) security_key_alloc(), ->key_alloc() (*) security_key_permission(), ->key_permission() Changed. These now take cred pointers rather than task pointers to refer to the security context. (4) sys_capset(). This has been simplified and uses less locking. The LSM functions it calls have been merged. (5) reparent_to_kthreadd(). This gives the current thread the same credentials as init by simply using commit_thread() to point that way. (6) __sigqueue_alloc() and switch_uid() __sigqueue_alloc() can't stop the target task from changing its creds beneath it, so this function gets a reference to the currently applicable user_struct which it then passes into the sigqueue struct it returns if successful. switch_uid() is now called from commit_creds(), and possibly should be folded into that. commit_creds() should take care of protecting __sigqueue_alloc(). (7) [sg]et[ug]id() and co and [sg]et_current_groups. The set functions now all use prepare_creds(), commit_creds() and abort_creds() to build and check a new set of credentials before applying it. security_task_set[ug]id() is called inside the prepared section. This guarantees that nothing else will affect the creds until we've finished. The calling of set_dumpable() has been moved into commit_creds(). Much of the functionality of set_user() has been moved into commit_creds(). The get functions all simply access the data directly. (8) security_task_prctl() and cap_task_prctl(). security_task_prctl() has been modified to return -ENOSYS if it doesn't want to handle a function, or otherwise return the return value directly rather than through an argument. Additionally, cap_task_prctl() now prepares a new set of credentials, even if it doesn't end up using it. (9) Keyrings. A number of changes have been made to the keyrings code: (a) switch_uid_keyring(), copy_keys(), exit_keys() and suid_keys() have all been dropped and built in to the credentials functions directly. They may want separating out again later. (b) key_alloc() and search_process_keyrings() now take a cred pointer rather than a task pointer to specify the security context. (c) copy_creds() gives a new thread within the same thread group a new thread keyring if its parent had one, otherwise it discards the thread keyring. (d) The authorisation key now points directly to the credentials to extend the search into rather pointing to the task that carries them. (e) Installing thread, process or session keyrings causes a new set of credentials to be created, even though it's not strictly necessary for process or session keyrings (they're shared). (10) Usermode helper. The usermode helper code now carries a cred struct pointer in its subprocess_info struct instead of a new session keyring pointer. This set of credentials is derived from init_cred and installed on the new process after it has been cloned. call_usermodehelper_setup() allocates the new credentials and call_usermodehelper_freeinfo() discards them if they haven't been used. A special cred function (prepare_usermodeinfo_creds()) is provided specifically for call_usermodehelper_setup() to call. call_usermodehelper_setkeys() adjusts the credentials to sport the supplied keyring as the new session keyring. (11) SELinux. SELinux has a number of changes, in addition to those to support the LSM interface changes mentioned above: (a) selinux_setprocattr() no longer does its check for whether the current ptracer can access processes with the new SID inside the lock that covers getting the ptracer's SID. Whilst this lock ensures that the check is done with the ptracer pinned, the result is only valid until the lock is released, so there's no point doing it inside the lock. (12) is_single_threaded(). This function has been extracted from selinux_setprocattr() and put into a file of its own in the lib/ directory as join_session_keyring() now wants to use it too. The code in SELinux just checked to see whether a task shared mm_structs with other tasks (CLONE_VM), but that isn't good enough. We really want to know if they're part of the same thread group (CLONE_THREAD). (13) nfsd. The NFS server daemon now has to use the COW credentials to set the credentials it is going to use. It really needs to pass the credentials down to the functions it calls, but it can't do that until other patches in this series have been applied. Signed-off-by: David Howells <dhowells@redhat.com> Acked-by: James Morris <jmorris@namei.org> Signed-off-by: James Morris <jmorris@namei.org>
2008-11-13 23:39:23 +00:00
goto bad_fork_cleanup_mm;
retval = copy_io(clone_flags, p);
if (retval)
goto bad_fork_cleanup_namespaces;
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
retval = copy_thread_tls(clone_flags, stack_start, stack_size, p, tls);
if (retval)
goto bad_fork_cleanup_io;
if (pid != &init_struct_pid) {
pid = alloc_pid(p->nsproxy->pid_ns_for_children);
fork: report pid reservation failure properly copy_process will report any failure in alloc_pid as ENOMEM currently which is misleading because the pid allocation might fail not only when the memory is short but also when the pid space is consumed already. The current man page even mentions this case: : EAGAIN : : A system-imposed limit on the number of threads was encountered. : There are a number of limits that may trigger this error: the : RLIMIT_NPROC soft resource limit (set via setrlimit(2)), which : limits the number of processes and threads for a real user ID, was : reached; the kernel's system-wide limit on the number of processes : and threads, /proc/sys/kernel/threads-max, was reached (see : proc(5)); or the maximum number of PIDs, /proc/sys/kernel/pid_max, : was reached (see proc(5)). so the current behavior is also incorrect wrt. documentation. POSIX man page also suggest returing EAGAIN when the process count limit is reached. This patch simply propagates error code from alloc_pid and makes sure we return -EAGAIN due to reservation failure. This will make behavior of fork closer to both our documentation and POSIX. alloc_pid might alsoo fail when the reaper in the pid namespace is dead (the namespace basically disallows all new processes) and there is no good error code which would match documented ones. We have traditionally returned ENOMEM for this case which is misleading as well but as per Eric W. Biederman this behavior is documented in man pid_namespaces(7) : If the "init" process of a PID namespace terminates, the kernel : terminates all of the processes in the namespace via a SIGKILL signal. : This behavior reflects the fact that the "init" process is essential for : the correct operation of a PID namespace. In this case, a subsequent : fork(2) into this PID namespace will fail with the error ENOMEM; it is : not possible to create a new processes in a PID namespace whose "init" : process has terminated. and introducing a new error code would be too risky so let's stick to ENOMEM for this case. Signed-off-by: Michal Hocko <mhocko@suse.cz> Cc: Oleg Nesterov <oleg@redhat.com> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-04-16 19:47:38 +00:00
if (IS_ERR(pid)) {
retval = PTR_ERR(pid);
fork: free thread in copy_process on failure When using this program (as root): #include <err.h> #include <stdio.h> #include <stdlib.h> #include <unistd.h> #include <sys/io.h> #include <sys/types.h> #include <sys/wait.h> #define ITER 1000 #define FORKERS 15 #define THREADS (6000/FORKERS) // 1850 is proc max static void fork_100_wait() { unsigned a, to_wait = 0; printf("\t%d forking %d\n", THREADS, getpid()); for (a = 0; a < THREADS; a++) { switch (fork()) { case 0: usleep(1000); exit(0); break; case -1: break; default: to_wait++; break; } } printf("\t%d forked from %d, waiting for %d\n", THREADS, getpid(), to_wait); for (a = 0; a < to_wait; a++) wait(NULL); printf("\t%d waited from %d\n", THREADS, getpid()); } static void run_forkers() { pid_t forkers[FORKERS]; unsigned a; for (a = 0; a < FORKERS; a++) { switch ((forkers[a] = fork())) { case 0: fork_100_wait(); exit(0); break; case -1: err(1, "DIE fork of %d'th forker", a); break; default: break; } } for (a = 0; a < FORKERS; a++) waitpid(forkers[a], NULL, 0); } int main() { unsigned a; int ret; ret = ioperm(10, 20, 0); if (ret < 0) err(1, "ioperm"); for (a = 0; a < ITER; a++) run_forkers(); return 0; } kmemleak reports many occurences of this leak: unreferenced object 0xffff8805917c8000 (size 8192): comm "fork-leak", pid 2932, jiffies 4295354292 (age 1871.028s) hex dump (first 32 bytes): ff ff ff ff ff ff ff ff ff ff ff ff ff ff ff ff ................ ff ff ff ff ff ff ff ff ff ff ff ff ff ff ff ff ................ backtrace: [<ffffffff814cfbf5>] kmemdup+0x25/0x50 [<ffffffff8103ab43>] copy_thread_tls+0x6c3/0x9a0 [<ffffffff81150174>] copy_process+0x1a84/0x5790 [<ffffffff811dc375>] wake_up_new_task+0x2d5/0x6f0 [<ffffffff8115411d>] _do_fork+0x12d/0x820 ... Due to the leakage of the memory items which should have been freed in arch/x86/kernel/process.c:exit_thread(). Make sure the memory is freed when fork fails later in copy_process. This is done by calling exit_thread with the thread to kill. Signed-off-by: Jiri Slaby <jslaby@suse.cz> Cc: "David S. Miller" <davem@davemloft.net> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: "James E.J. Bottomley" <jejb@parisc-linux.org> Cc: Aurelien Jacquiot <a-jacquiot@ti.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Chen Liqin <liqin.linux@gmail.com> Cc: Chris Metcalf <cmetcalf@mellanox.com> Cc: Chris Zankel <chris@zankel.net> Cc: David Howells <dhowells@redhat.com> Cc: Fenghua Yu <fenghua.yu@intel.com> Cc: Geert Uytterhoeven <geert@linux-m68k.org> Cc: Guan Xuetao <gxt@mprc.pku.edu.cn> Cc: Haavard Skinnemoen <hskinnemoen@gmail.com> Cc: Hans-Christian Egtvedt <egtvedt@samfundet.no> Cc: Heiko Carstens <heiko.carstens@de.ibm.com> Cc: Helge Deller <deller@gmx.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: Ivan Kokshaysky <ink@jurassic.park.msu.ru> Cc: James Hogan <james.hogan@imgtec.com> Cc: Jeff Dike <jdike@addtoit.com> Cc: Jesper Nilsson <jesper.nilsson@axis.com> Cc: Jiri Slaby <jslaby@suse.cz> Cc: Jonas Bonn <jonas@southpole.se> Cc: Koichi Yasutake <yasutake.koichi@jp.panasonic.com> Cc: Lennox Wu <lennox.wu@gmail.com> Cc: Ley Foon Tan <lftan@altera.com> Cc: Mark Salter <msalter@redhat.com> Cc: Martin Schwidefsky <schwidefsky@de.ibm.com> Cc: Matt Turner <mattst88@gmail.com> Cc: Max Filippov <jcmvbkbc@gmail.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Michal Simek <monstr@monstr.eu> Cc: Mikael Starvik <starvik@axis.com> Cc: Paul Mackerras <paulus@samba.org> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Ralf Baechle <ralf@linux-mips.org> Cc: Rich Felker <dalias@libc.org> Cc: Richard Henderson <rth@twiddle.net> Cc: Richard Kuo <rkuo@codeaurora.org> Cc: Richard Weinberger <richard@nod.at> Cc: Russell King <linux@arm.linux.org.uk> Cc: Steven Miao <realmz6@gmail.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Tony Luck <tony.luck@intel.com> Cc: Vineet Gupta <vgupta@synopsys.com> Cc: Will Deacon <will.deacon@arm.com> Cc: Yoshinori Sato <ysato@users.sourceforge.jp> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-05-21 00:00:25 +00:00
goto bad_fork_cleanup_thread;
fork: report pid reservation failure properly copy_process will report any failure in alloc_pid as ENOMEM currently which is misleading because the pid allocation might fail not only when the memory is short but also when the pid space is consumed already. The current man page even mentions this case: : EAGAIN : : A system-imposed limit on the number of threads was encountered. : There are a number of limits that may trigger this error: the : RLIMIT_NPROC soft resource limit (set via setrlimit(2)), which : limits the number of processes and threads for a real user ID, was : reached; the kernel's system-wide limit on the number of processes : and threads, /proc/sys/kernel/threads-max, was reached (see : proc(5)); or the maximum number of PIDs, /proc/sys/kernel/pid_max, : was reached (see proc(5)). so the current behavior is also incorrect wrt. documentation. POSIX man page also suggest returing EAGAIN when the process count limit is reached. This patch simply propagates error code from alloc_pid and makes sure we return -EAGAIN due to reservation failure. This will make behavior of fork closer to both our documentation and POSIX. alloc_pid might alsoo fail when the reaper in the pid namespace is dead (the namespace basically disallows all new processes) and there is no good error code which would match documented ones. We have traditionally returned ENOMEM for this case which is misleading as well but as per Eric W. Biederman this behavior is documented in man pid_namespaces(7) : If the "init" process of a PID namespace terminates, the kernel : terminates all of the processes in the namespace via a SIGKILL signal. : This behavior reflects the fact that the "init" process is essential for : the correct operation of a PID namespace. In this case, a subsequent : fork(2) into this PID namespace will fail with the error ENOMEM; it is : not possible to create a new processes in a PID namespace whose "init" : process has terminated. and introducing a new error code would be too risky so let's stick to ENOMEM for this case. Signed-off-by: Michal Hocko <mhocko@suse.cz> Cc: Oleg Nesterov <oleg@redhat.com> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-04-16 19:47:38 +00:00
}
}
p->set_child_tid = (clone_flags & CLONE_CHILD_SETTID) ? child_tidptr : NULL;
/*
* Clear TID on mm_release()?
*/
p->clear_child_tid = (clone_flags & CLONE_CHILD_CLEARTID) ? child_tidptr : NULL;
#ifdef CONFIG_BLOCK
p->plug = NULL;
#endif
#ifdef CONFIG_FUTEX
p->robust_list = NULL;
#ifdef CONFIG_COMPAT
p->compat_robust_list = NULL;
#endif
INIT_LIST_HEAD(&p->pi_state_list);
p->pi_state_cache = NULL;
#endif
[PATCH] Fix sigaltstack corruption among cloned threads This patch fixes alternate signal stack corruption among cloned threads with CLONE_SIGHAND (and CLONE_VM) for linux-2.6.16-rc6. The value of alternate signal stack is currently inherited after a call of clone(... CLONE_SIGHAND | CLONE_VM). But if sigaltstack is set by a parent thread, and then if multiple cloned child threads (+ parent threads) call signal handler at the same time, some threads may be conflicted - because they share to use the same alternative signal stack region. Finally they get sigsegv. It's an undesirable race condition. Note that child threads created from NPTL pthread_create() also hit this conflict when the parent thread uses sigaltstack, without my patch. To fix this problem, this patch clears the child threads' sigaltstack information like exec(). This behavior follows the SUSv3 specification. In SUSv3, pthread_create() says "The alternate stack shall not be inherited (when new threads are initialized)". It means that sigaltstack should be cleared when sigaltstack memory space is shared by cloned threads with CLONE_SIGHAND. Note that I chose "if (clone_flags & CLONE_SIGHAND)" line because: - If clone_flags line is not existed, fork() does not inherit sigaltstack. - CLONE_VM is another choice, but vfork() does not inherit sigaltstack. - CLONE_SIGHAND implies CLONE_VM, and it looks suitable. - CLONE_THREAD is another candidate, and includes CLONE_SIGHAND + CLONE_VM, but this flag has a bit different semantics. I decided to use CLONE_SIGHAND. [ Changed to test for CLONE_VM && !CLONE_VFORK after discussion --Linus ] Signed-off-by: GOTO Masanori <gotom@sanori.org> Cc: Roland McGrath <roland@redhat.com> Cc: Ingo Molnar <mingo@elte.hu> Acked-by: Linus Torvalds <torvalds@osdl.org> Cc: Ulrich Drepper <drepper@redhat.com> Cc: Jakub Jelinek <jakub@redhat.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-03-14 05:20:44 +00:00
/*
* sigaltstack should be cleared when sharing the same VM
*/
if ((clone_flags & (CLONE_VM|CLONE_VFORK)) == CLONE_VM)
signals/sigaltstack: Implement SS_AUTODISARM flag This patch implements the SS_AUTODISARM flag that can be OR-ed with SS_ONSTACK when forming ss_flags. When this flag is set, sigaltstack will be disabled when entering the signal handler; more precisely, after saving sas to uc_stack. When leaving the signal handler, the sigaltstack is restored by uc_stack. When this flag is used, it is safe to switch from sighandler with swapcontext(). Without this flag, the subsequent signal will corrupt the state of the switched-away sighandler. To detect the support of this functionality, one can do: err = sigaltstack(SS_DISABLE | SS_AUTODISARM); if (err && errno == EINVAL) unsupported(); Signed-off-by: Stas Sergeev <stsp@list.ru> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Aleksa Sarai <cyphar@cyphar.com> Cc: Amanieu d'Antras <amanieu@gmail.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Andy Lutomirski <luto@amacapital.net> Cc: Borislav Petkov <bp@alien8.de> Cc: Brian Gerst <brgerst@gmail.com> Cc: Denys Vlasenko <dvlasenk@redhat.com> Cc: Eric W. Biederman <ebiederm@xmission.com> Cc: Frederic Weisbecker <fweisbec@gmail.com> Cc: H. Peter Anvin <hpa@zytor.com> Cc: Heinrich Schuchardt <xypron.glpk@gmx.de> Cc: Jason Low <jason.low2@hp.com> Cc: Josh Triplett <josh@joshtriplett.org> Cc: Konstantin Khlebnikov <khlebnikov@yandex-team.ru> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Palmer Dabbelt <palmer@dabbelt.com> Cc: Paul Moore <pmoore@redhat.com> Cc: Pavel Emelyanov <xemul@parallels.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Richard Weinberger <richard@nod.at> Cc: Sasha Levin <sasha.levin@oracle.com> Cc: Shuah Khan <shuahkh@osg.samsung.com> Cc: Tejun Heo <tj@kernel.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Vladimir Davydov <vdavydov@parallels.com> Cc: linux-api@vger.kernel.org Cc: linux-kernel@vger.kernel.org Link: http://lkml.kernel.org/r/1460665206-13646-4-git-send-email-stsp@list.ru Signed-off-by: Ingo Molnar <mingo@kernel.org>
2016-04-14 20:20:04 +00:00
sas_ss_reset(p);
[PATCH] Fix sigaltstack corruption among cloned threads This patch fixes alternate signal stack corruption among cloned threads with CLONE_SIGHAND (and CLONE_VM) for linux-2.6.16-rc6. The value of alternate signal stack is currently inherited after a call of clone(... CLONE_SIGHAND | CLONE_VM). But if sigaltstack is set by a parent thread, and then if multiple cloned child threads (+ parent threads) call signal handler at the same time, some threads may be conflicted - because they share to use the same alternative signal stack region. Finally they get sigsegv. It's an undesirable race condition. Note that child threads created from NPTL pthread_create() also hit this conflict when the parent thread uses sigaltstack, without my patch. To fix this problem, this patch clears the child threads' sigaltstack information like exec(). This behavior follows the SUSv3 specification. In SUSv3, pthread_create() says "The alternate stack shall not be inherited (when new threads are initialized)". It means that sigaltstack should be cleared when sigaltstack memory space is shared by cloned threads with CLONE_SIGHAND. Note that I chose "if (clone_flags & CLONE_SIGHAND)" line because: - If clone_flags line is not existed, fork() does not inherit sigaltstack. - CLONE_VM is another choice, but vfork() does not inherit sigaltstack. - CLONE_SIGHAND implies CLONE_VM, and it looks suitable. - CLONE_THREAD is another candidate, and includes CLONE_SIGHAND + CLONE_VM, but this flag has a bit different semantics. I decided to use CLONE_SIGHAND. [ Changed to test for CLONE_VM && !CLONE_VFORK after discussion --Linus ] Signed-off-by: GOTO Masanori <gotom@sanori.org> Cc: Roland McGrath <roland@redhat.com> Cc: Ingo Molnar <mingo@elte.hu> Acked-by: Linus Torvalds <torvalds@osdl.org> Cc: Ulrich Drepper <drepper@redhat.com> Cc: Jakub Jelinek <jakub@redhat.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-03-14 05:20:44 +00:00
/*
* Syscall tracing and stepping should be turned off in the
* child regardless of CLONE_PTRACE.
*/
user_disable_single_step(p);
clear_tsk_thread_flag(p, TIF_SYSCALL_TRACE);
[PATCH] UML Support - Ptrace: adds the host SYSEMU support, for UML and general usage Jeff Dike <jdike@addtoit.com>, Paolo 'Blaisorblade' Giarrusso <blaisorblade_spam@yahoo.it>, Bodo Stroesser <bstroesser@fujitsu-siemens.com> Adds a new ptrace(2) mode, called PTRACE_SYSEMU, resembling PTRACE_SYSCALL except that the kernel does not execute the requested syscall; this is useful to improve performance for virtual environments, like UML, which want to run the syscall on their own. In fact, using PTRACE_SYSCALL means stopping child execution twice, on entry and on exit, and each time you also have two context switches; with SYSEMU you avoid the 2nd stop and so save two context switches per syscall. Also, some architectures don't have support in the host for changing the syscall number via ptrace(), which is currently needed to skip syscall execution (UML turns any syscall into getpid() to avoid it being executed on the host). Fixing that is hard, while SYSEMU is easier to implement. * This version of the patch includes some suggestions of Jeff Dike to avoid adding any instructions to the syscall fast path, plus some other little changes, by myself, to make it work even when the syscall is executed with SYSENTER (but I'm unsure about them). It has been widely tested for quite a lot of time. * Various fixed were included to handle the various switches between various states, i.e. when for instance a syscall entry is traced with one of PT_SYSCALL / _SYSEMU / _SINGLESTEP and another one is used on exit. Basically, this is done by remembering which one of them was used even after the call to ptrace_notify(). * We're combining TIF_SYSCALL_EMU with TIF_SYSCALL_TRACE or TIF_SINGLESTEP to make do_syscall_trace() notice that the current syscall was started with SYSEMU on entry, so that no notification ought to be done in the exit path; this is a bit of a hack, so this problem is solved in another way in next patches. * Also, the effects of the patch: "Ptrace - i386: fix Syscall Audit interaction with singlestep" are cancelled; they are restored back in the last patch of this series. Detailed descriptions of the patches doing this kind of processing follow (but I've already summed everything up). * Fix behaviour when changing interception kind #1. In do_syscall_trace(), we check the status of the TIF_SYSCALL_EMU flag only after doing the debugger notification; but the debugger might have changed the status of this flag because he continued execution with PTRACE_SYSCALL, so this is wrong. This patch fixes it by saving the flag status before calling ptrace_notify(). * Fix behaviour when changing interception kind #2: avoid intercepting syscall on return when using SYSCALL again. A guest process switching from using PTRACE_SYSEMU to PTRACE_SYSCALL crashes. The problem is in arch/i386/kernel/entry.S. The current SYSEMU patch inhibits the syscall-handler to be called, but does not prevent do_syscall_trace() to be called after this for syscall completion interception. The appended patch fixes this. It reuses the flag TIF_SYSCALL_EMU to remember "we come from PTRACE_SYSEMU and now are in PTRACE_SYSCALL", since the flag is unused in the depicted situation. * Fix behaviour when changing interception kind #3: avoid intercepting syscall on return when using SINGLESTEP. When testing 2.6.9 and the skas3.v6 patch, with my latest patch and had problems with singlestepping on UML in SKAS with SYSEMU. It looped receiving SIGTRAPs without moving forward. EIP of the traced process was the same for all SIGTRAPs. What's missing is to handle switching from PTRACE_SYSCALL_EMU to PTRACE_SINGLESTEP in a way very similar to what is done for the change from PTRACE_SYSCALL_EMU to PTRACE_SYSCALL_TRACE. I.e., after calling ptrace(PTRACE_SYSEMU), on the return path, the debugger is notified and then wake ups the process; the syscall is executed (or skipped, when do_syscall_trace() returns 0, i.e. when using PTRACE_SYSEMU), and do_syscall_trace() is called again. Since we are on the return path of a SYSEMU'd syscall, if the wake up is performed through ptrace(PTRACE_SYSCALL), we must still avoid notifying the parent of the syscall exit. Now, this behaviour is extended even to resuming with PTRACE_SINGLESTEP. Signed-off-by: Paolo 'Blaisorblade' Giarrusso <blaisorblade@yahoo.it> Cc: Jeff Dike <jdike@addtoit.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-09-03 22:57:18 +00:00
#ifdef TIF_SYSCALL_EMU
clear_tsk_thread_flag(p, TIF_SYSCALL_EMU);
#endif
clear_all_latency_tracing(p);
/* ok, now we should be set up.. */
p->pid = pid_nr(pid);
if (clone_flags & CLONE_THREAD) {
p->exit_signal = -1;
p->group_leader = current->group_leader;
p->tgid = current->tgid;
} else {
if (clone_flags & CLONE_PARENT)
p->exit_signal = current->group_leader->exit_signal;
else
p->exit_signal = (clone_flags & CSIGNAL);
p->group_leader = p;
p->tgid = p->pid;
}
writeback: per task dirty rate limit Add two fields to task_struct. 1) account dirtied pages in the individual tasks, for accuracy 2) per-task balance_dirty_pages() call intervals, for flexibility The balance_dirty_pages() call interval (ie. nr_dirtied_pause) will scale near-sqrt to the safety gap between dirty pages and threshold. The main problem of per-task nr_dirtied is, if 1k+ tasks start dirtying pages at exactly the same time, each task will be assigned a large initial nr_dirtied_pause, so that the dirty threshold will be exceeded long before each task reached its nr_dirtied_pause and hence call balance_dirty_pages(). The solution is to watch for the number of pages dirtied on each CPU in between the calls into balance_dirty_pages(). If it exceeds ratelimit_pages (3% dirty threshold), force call balance_dirty_pages() for a chance to set bdi->dirty_exceeded. In normal situations, this safeguarding condition is not expected to trigger at all. On the sqrt in dirty_poll_interval(): It will serve as an initial guess when dirty pages are still in the freerun area. When dirty pages are floating inside the dirty control scope [freerun, limit], a followup patch will use some refined dirty poll interval to get the desired pause time. thresh-dirty (MB) sqrt 1 16 2 22 4 32 8 45 16 64 32 90 64 128 128 181 256 256 512 362 1024 512 The above table means, given 1MB (or 1GB) gap and the dd tasks polling balance_dirty_pages() on every 16 (or 512) pages, the dirty limit won't be exceeded as long as there are less than 16 (or 512) concurrent dd's. So sqrt naturally leads to less overheads and more safe concurrent tasks for large memory servers, which have large (thresh-freerun) gaps. peter: keep the per-CPU ratelimit for safeguarding the 1k+ tasks case CC: Peter Zijlstra <a.p.zijlstra@chello.nl> Reviewed-by: Andrea Righi <andrea@betterlinux.com> Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-06-12 00:10:12 +00:00
p->nr_dirtied = 0;
p->nr_dirtied_pause = 128 >> (PAGE_SHIFT - 10);
p->dirty_paused_when = 0;
writeback: per task dirty rate limit Add two fields to task_struct. 1) account dirtied pages in the individual tasks, for accuracy 2) per-task balance_dirty_pages() call intervals, for flexibility The balance_dirty_pages() call interval (ie. nr_dirtied_pause) will scale near-sqrt to the safety gap between dirty pages and threshold. The main problem of per-task nr_dirtied is, if 1k+ tasks start dirtying pages at exactly the same time, each task will be assigned a large initial nr_dirtied_pause, so that the dirty threshold will be exceeded long before each task reached its nr_dirtied_pause and hence call balance_dirty_pages(). The solution is to watch for the number of pages dirtied on each CPU in between the calls into balance_dirty_pages(). If it exceeds ratelimit_pages (3% dirty threshold), force call balance_dirty_pages() for a chance to set bdi->dirty_exceeded. In normal situations, this safeguarding condition is not expected to trigger at all. On the sqrt in dirty_poll_interval(): It will serve as an initial guess when dirty pages are still in the freerun area. When dirty pages are floating inside the dirty control scope [freerun, limit], a followup patch will use some refined dirty poll interval to get the desired pause time. thresh-dirty (MB) sqrt 1 16 2 22 4 32 8 45 16 64 32 90 64 128 128 181 256 256 512 362 1024 512 The above table means, given 1MB (or 1GB) gap and the dd tasks polling balance_dirty_pages() on every 16 (or 512) pages, the dirty limit won't be exceeded as long as there are less than 16 (or 512) concurrent dd's. So sqrt naturally leads to less overheads and more safe concurrent tasks for large memory servers, which have large (thresh-freerun) gaps. peter: keep the per-CPU ratelimit for safeguarding the 1k+ tasks case CC: Peter Zijlstra <a.p.zijlstra@chello.nl> Reviewed-by: Andrea Righi <andrea@betterlinux.com> Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-06-12 00:10:12 +00:00
p->pdeath_signal = 0;
INIT_LIST_HEAD(&p->thread_group);
p->task_works = NULL;
cgroup_threadgroup_change_begin(current);
/*
* Ensure that the cgroup subsystem policies allow the new process to be
* forked. It should be noted the the new process's css_set can be changed
* between here and cgroup_post_fork() if an organisation operation is in
* progress.
*/
retval = cgroup_can_fork(p);
if (retval)
goto bad_fork_free_pid;
/*
* Make it visible to the rest of the system, but dont wake it up yet.
* Need tasklist lock for parent etc handling!
*/
write_lock_irq(&tasklist_lock);
/* CLONE_PARENT re-uses the old parent */
if (clone_flags & (CLONE_PARENT|CLONE_THREAD)) {
p->real_parent = current->real_parent;
p->parent_exec_id = current->parent_exec_id;
} else {
p->real_parent = current;
p->parent_exec_id = current->self_exec_id;
}
spin_lock(&current->sighand->siglock);
/*
* Copy seccomp details explicitly here, in case they were changed
* before holding sighand lock.
*/
copy_seccomp(p);
/*
* Process group and session signals need to be delivered to just the
* parent before the fork or both the parent and the child after the
* fork. Restart if a signal comes in before we add the new process to
* it's process group.
* A fatal signal pending means that current will exit, so the new
* thread can't slip out of an OOM kill (or normal SIGKILL).
*/
recalc_sigpending();
if (signal_pending(current)) {
spin_unlock(&current->sighand->siglock);
write_unlock_irq(&tasklist_lock);
retval = -ERESTARTNOINTR;
goto bad_fork_cancel_cgroup;
}
if (likely(p->pid)) {
ptrace_init_task(p, (clone_flags & CLONE_PTRACE) || trace);
init_task_pid(p, PIDTYPE_PID, pid);
if (thread_group_leader(p)) {
init_task_pid(p, PIDTYPE_PGID, task_pgrp(current));
init_task_pid(p, PIDTYPE_SID, task_session(current));
if (is_child_reaper(pid)) {
ns_of_pid(pid)->child_reaper = p;
p->signal->flags |= SIGNAL_UNKILLABLE;
}
p->signal->leader_pid = pid;
p->signal->tty = tty_kref_get(current->signal->tty);
prctl: propagate has_child_subreaper flag to every descendant If process forks some children when it has is_child_subreaper flag enabled they will inherit has_child_subreaper flag - first group, when is_child_subreaper is disabled forked children will not inherit it - second group. So child-subreaper does not reparent all his descendants when their parents die. Having these two differently behaving groups can lead to confusion. Also it is a problem for CRIU, as when we restore process tree we need to somehow determine which descendants belong to which group and much harder - to put them exactly to these group. To simplify these we can add a propagation of has_child_subreaper flag on PR_SET_CHILD_SUBREAPER, walking all descendants of child- subreaper to setup has_child_subreaper flag. In common cases when process like systemd first sets itself to be a child-subreaper and only after that forks its services, we will have zero-length list of descendants to walk. Testing with binary subtree of 2^15 processes prctl took < 0.007 sec and has shown close to linear dependency(~0.2 * n * usec) on lower numbers of processes. Moreover, I doubt someone intentionaly pre-forks the children whitch should reparent to init before becoming subreaper, because some our ancestor migh have had is_child_subreaper flag while forking our sub-tree and our childs will all inherit has_child_subreaper flag, and we have no way to influence it. And only way to check if we have no has_child_subreaper flag is to create some childs, kill them and see where they will reparent to. Using walk_process_tree helper to walk subtree, thanks to Oleg! Timing seems to be the same. Optimize: a) When descendant already has has_child_subreaper flag all his subtree has it too already. * for a) to be true need to move has_child_subreaper inheritance under the same tasklist_lock with adding task to its ->real_parent->children as without it process can inherit zero has_child_subreaper, then we set 1 to it's parent flag, check that parent has no more children, and only after child with wrong flag is added to the tree. * Also make these inheritance more clear by using real_parent instead of current, as on clone(CLONE_PARENT) if current has is_child_subreaper and real_parent has no is_child_subreaper or has_child_subreaper, child will have has_child_subreaper flag set without actually having a subreaper in it's ancestors. b) When some descendant is child_reaper, it's subtree is in different pidns from us(original child-subreaper) and processes from other pidns will never reparent to us. So we can skip their(a,b) subtree from walk. v2: switch to walk_process_tree() general helper, move has_child_subreaper inheritance v3: remove csr_descendant leftover, change current to real_parent in has_child_subreaper inheritance v4: small commit message fix Fixes: ebec18a6d3aa ("prctl: add PR_{SET,GET}_CHILD_SUBREAPER to allow simple process supervision") Signed-off-by: Pavel Tikhomirov <ptikhomirov@virtuozzo.com> Reviewed-by: Oleg Nesterov <oleg@redhat.com> Signed-off-by: Eric W. Biederman <ebiederm@xmission.com>
2017-01-30 15:06:12 +00:00
/*
* Inherit has_child_subreaper flag under the same
* tasklist_lock with adding child to the process tree
* for propagate_has_child_subreaper optimization.
*/
p->signal->has_child_subreaper = p->real_parent->signal->has_child_subreaper ||
p->real_parent->signal->is_child_subreaper;
list_add_tail(&p->sibling, &p->real_parent->children);
list_add_tail_rcu(&p->tasks, &init_task.tasks);
attach_pid(p, PIDTYPE_PGID);
attach_pid(p, PIDTYPE_SID);
__this_cpu_inc(process_counts);
} else {
current->signal->nr_threads++;
atomic_inc(&current->signal->live);
atomic_inc(&current->signal->sigcnt);
list_add_tail_rcu(&p->thread_group,
&p->group_leader->thread_group);
introduce for_each_thread() to replace the buggy while_each_thread() while_each_thread() and next_thread() should die, almost every lockless usage is wrong. 1. Unless g == current, the lockless while_each_thread() is not safe. while_each_thread(g, t) can loop forever if g exits, next_thread() can't reach the unhashed thread in this case. Note that this can happen even if g is the group leader, it can exec. 2. Even if while_each_thread() itself was correct, people often use it wrongly. It was never safe to just take rcu_read_lock() and loop unless you verify that pid_alive(g) == T, even the first next_thread() can point to the already freed/reused memory. This patch adds signal_struct->thread_head and task->thread_node to create the normal rcu-safe list with the stable head. The new for_each_thread(g, t) helper is always safe under rcu_read_lock() as long as this task_struct can't go away. Note: of course it is ugly to have both task_struct->thread_node and the old task_struct->thread_group, we will kill it later, after we change the users of while_each_thread() to use for_each_thread(). Perhaps we can kill it even before we convert all users, we can reimplement next_thread(t) using the new thread_head/thread_node. But we can't do this right now because this will lead to subtle behavioural changes. For example, do/while_each_thread() always sees at least one task, while for_each_thread() can do nothing if the whole thread group has died. Or thread_group_empty(), currently its semantics is not clear unless thread_group_leader(p) and we need to audit the callers before we can change it. So this patch adds the new interface which has to coexist with the old one for some time, hopefully the next changes will be more or less straightforward and the old one will go away soon. Signed-off-by: Oleg Nesterov <oleg@redhat.com> Reviewed-by: Sergey Dyasly <dserrg@gmail.com> Tested-by: Sergey Dyasly <dserrg@gmail.com> Reviewed-by: Sameer Nanda <snanda@chromium.org> Acked-by: David Rientjes <rientjes@google.com> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: Frederic Weisbecker <fweisbec@gmail.com> Cc: Mandeep Singh Baines <msb@chromium.org> Cc: "Ma, Xindong" <xindong.ma@intel.com> Cc: Michal Hocko <mhocko@suse.cz> Cc: "Tu, Xiaobing" <xiaobing.tu@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-01-21 23:49:56 +00:00
list_add_tail_rcu(&p->thread_node,
&p->signal->thread_head);
}
attach_pid(p, PIDTYPE_PID);
nr_threads++;
}
total_forks++;
spin_unlock(&current->sighand->siglock);
syscall_tracepoint_update(p);
write_unlock_irq(&tasklist_lock);
proc_fork_connector(p);
cgroup_post_fork(p);
cgroup_threadgroup_change_end(current);
perf: Do the big rename: Performance Counters -> Performance Events Bye-bye Performance Counters, welcome Performance Events! In the past few months the perfcounters subsystem has grown out its initial role of counting hardware events, and has become (and is becoming) a much broader generic event enumeration, reporting, logging, monitoring, analysis facility. Naming its core object 'perf_counter' and naming the subsystem 'perfcounters' has become more and more of a misnomer. With pending code like hw-breakpoints support the 'counter' name is less and less appropriate. All in one, we've decided to rename the subsystem to 'performance events' and to propagate this rename through all fields, variables and API names. (in an ABI compatible fashion) The word 'event' is also a bit shorter than 'counter' - which makes it slightly more convenient to write/handle as well. Thanks goes to Stephane Eranian who first observed this misnomer and suggested a rename. User-space tooling and ABI compatibility is not affected - this patch should be function-invariant. (Also, defconfigs were not touched to keep the size down.) This patch has been generated via the following script: FILES=$(find * -type f | grep -vE 'oprofile|[^K]config') sed -i \ -e 's/PERF_EVENT_/PERF_RECORD_/g' \ -e 's/PERF_COUNTER/PERF_EVENT/g' \ -e 's/perf_counter/perf_event/g' \ -e 's/nb_counters/nb_events/g' \ -e 's/swcounter/swevent/g' \ -e 's/tpcounter_event/tp_event/g' \ $FILES for N in $(find . -name perf_counter.[ch]); do M=$(echo $N | sed 's/perf_counter/perf_event/g') mv $N $M done FILES=$(find . -name perf_event.*) sed -i \ -e 's/COUNTER_MASK/REG_MASK/g' \ -e 's/COUNTER/EVENT/g' \ -e 's/\<event\>/event_id/g' \ -e 's/counter/event/g' \ -e 's/Counter/Event/g' \ $FILES ... to keep it as correct as possible. This script can also be used by anyone who has pending perfcounters patches - it converts a Linux kernel tree over to the new naming. We tried to time this change to the point in time where the amount of pending patches is the smallest: the end of the merge window. Namespace clashes were fixed up in a preparatory patch - and some stylistic fallout will be fixed up in a subsequent patch. ( NOTE: 'counters' are still the proper terminology when we deal with hardware registers - and these sed scripts are a bit over-eager in renaming them. I've undone some of that, but in case there's something left where 'counter' would be better than 'event' we can undo that on an individual basis instead of touching an otherwise nicely automated patch. ) Suggested-by: Stephane Eranian <eranian@google.com> Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl> Acked-by: Paul Mackerras <paulus@samba.org> Reviewed-by: Arjan van de Ven <arjan@linux.intel.com> Cc: Mike Galbraith <efault@gmx.de> Cc: Arnaldo Carvalho de Melo <acme@redhat.com> Cc: Frederic Weisbecker <fweisbec@gmail.com> Cc: Steven Rostedt <rostedt@goodmis.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: David Howells <dhowells@redhat.com> Cc: Kyle McMartin <kyle@mcmartin.ca> Cc: Martin Schwidefsky <schwidefsky@de.ibm.com> Cc: "David S. Miller" <davem@davemloft.net> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: <linux-arch@vger.kernel.org> LKML-Reference: <new-submission> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-09-21 10:02:48 +00:00
perf_event_fork(p);
tracepoint: add tracepoints for debugging oom_score_adj oom_score_adj is used for guarding processes from OOM-Killer. One of problem is that it's inherited at fork(). When a daemon set oom_score_adj and make children, it's hard to know where the value is set. This patch adds some tracepoints useful for debugging. This patch adds 3 trace points. - creating new task - renaming a task (exec) - set oom_score_adj To debug, users need to enable some trace pointer. Maybe filtering is useful as # EVENT=/sys/kernel/debug/tracing/events/task/ # echo "oom_score_adj != 0" > $EVENT/task_newtask/filter # echo "oom_score_adj != 0" > $EVENT/task_rename/filter # echo 1 > $EVENT/enable # EVENT=/sys/kernel/debug/tracing/events/oom/ # echo 1 > $EVENT/enable output will be like this. # grep oom /sys/kernel/debug/tracing/trace bash-7699 [007] d..3 5140.744510: oom_score_adj_update: pid=7699 comm=bash oom_score_adj=-1000 bash-7699 [007] ...1 5151.818022: task_newtask: pid=7729 comm=bash clone_flags=1200011 oom_score_adj=-1000 ls-7729 [003] ...2 5151.818504: task_rename: pid=7729 oldcomm=bash newcomm=ls oom_score_adj=-1000 bash-7699 [002] ...1 5175.701468: task_newtask: pid=7730 comm=bash clone_flags=1200011 oom_score_adj=-1000 grep-7730 [007] ...2 5175.701993: task_rename: pid=7730 oldcomm=bash newcomm=grep oom_score_adj=-1000 Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Acked-by: David Rientjes <rientjes@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-01-10 23:08:09 +00:00
trace_task_newtask(p, clone_flags);
uprobe_copy_process(p, clone_flags);
tracepoint: add tracepoints for debugging oom_score_adj oom_score_adj is used for guarding processes from OOM-Killer. One of problem is that it's inherited at fork(). When a daemon set oom_score_adj and make children, it's hard to know where the value is set. This patch adds some tracepoints useful for debugging. This patch adds 3 trace points. - creating new task - renaming a task (exec) - set oom_score_adj To debug, users need to enable some trace pointer. Maybe filtering is useful as # EVENT=/sys/kernel/debug/tracing/events/task/ # echo "oom_score_adj != 0" > $EVENT/task_newtask/filter # echo "oom_score_adj != 0" > $EVENT/task_rename/filter # echo 1 > $EVENT/enable # EVENT=/sys/kernel/debug/tracing/events/oom/ # echo 1 > $EVENT/enable output will be like this. # grep oom /sys/kernel/debug/tracing/trace bash-7699 [007] d..3 5140.744510: oom_score_adj_update: pid=7699 comm=bash oom_score_adj=-1000 bash-7699 [007] ...1 5151.818022: task_newtask: pid=7729 comm=bash clone_flags=1200011 oom_score_adj=-1000 ls-7729 [003] ...2 5151.818504: task_rename: pid=7729 oldcomm=bash newcomm=ls oom_score_adj=-1000 bash-7699 [002] ...1 5175.701468: task_newtask: pid=7730 comm=bash clone_flags=1200011 oom_score_adj=-1000 grep-7730 [007] ...2 5175.701993: task_rename: pid=7730 oldcomm=bash newcomm=grep oom_score_adj=-1000 Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Acked-by: David Rientjes <rientjes@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-01-10 23:08:09 +00:00
return p;
bad_fork_cancel_cgroup:
cgroup_cancel_fork(p);
bad_fork_free_pid:
cgroup_threadgroup_change_end(current);
if (pid != &init_struct_pid)
free_pid(pid);
fork: free thread in copy_process on failure When using this program (as root): #include <err.h> #include <stdio.h> #include <stdlib.h> #include <unistd.h> #include <sys/io.h> #include <sys/types.h> #include <sys/wait.h> #define ITER 1000 #define FORKERS 15 #define THREADS (6000/FORKERS) // 1850 is proc max static void fork_100_wait() { unsigned a, to_wait = 0; printf("\t%d forking %d\n", THREADS, getpid()); for (a = 0; a < THREADS; a++) { switch (fork()) { case 0: usleep(1000); exit(0); break; case -1: break; default: to_wait++; break; } } printf("\t%d forked from %d, waiting for %d\n", THREADS, getpid(), to_wait); for (a = 0; a < to_wait; a++) wait(NULL); printf("\t%d waited from %d\n", THREADS, getpid()); } static void run_forkers() { pid_t forkers[FORKERS]; unsigned a; for (a = 0; a < FORKERS; a++) { switch ((forkers[a] = fork())) { case 0: fork_100_wait(); exit(0); break; case -1: err(1, "DIE fork of %d'th forker", a); break; default: break; } } for (a = 0; a < FORKERS; a++) waitpid(forkers[a], NULL, 0); } int main() { unsigned a; int ret; ret = ioperm(10, 20, 0); if (ret < 0) err(1, "ioperm"); for (a = 0; a < ITER; a++) run_forkers(); return 0; } kmemleak reports many occurences of this leak: unreferenced object 0xffff8805917c8000 (size 8192): comm "fork-leak", pid 2932, jiffies 4295354292 (age 1871.028s) hex dump (first 32 bytes): ff ff ff ff ff ff ff ff ff ff ff ff ff ff ff ff ................ ff ff ff ff ff ff ff ff ff ff ff ff ff ff ff ff ................ backtrace: [<ffffffff814cfbf5>] kmemdup+0x25/0x50 [<ffffffff8103ab43>] copy_thread_tls+0x6c3/0x9a0 [<ffffffff81150174>] copy_process+0x1a84/0x5790 [<ffffffff811dc375>] wake_up_new_task+0x2d5/0x6f0 [<ffffffff8115411d>] _do_fork+0x12d/0x820 ... Due to the leakage of the memory items which should have been freed in arch/x86/kernel/process.c:exit_thread(). Make sure the memory is freed when fork fails later in copy_process. This is done by calling exit_thread with the thread to kill. Signed-off-by: Jiri Slaby <jslaby@suse.cz> Cc: "David S. Miller" <davem@davemloft.net> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: "James E.J. Bottomley" <jejb@parisc-linux.org> Cc: Aurelien Jacquiot <a-jacquiot@ti.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Chen Liqin <liqin.linux@gmail.com> Cc: Chris Metcalf <cmetcalf@mellanox.com> Cc: Chris Zankel <chris@zankel.net> Cc: David Howells <dhowells@redhat.com> Cc: Fenghua Yu <fenghua.yu@intel.com> Cc: Geert Uytterhoeven <geert@linux-m68k.org> Cc: Guan Xuetao <gxt@mprc.pku.edu.cn> Cc: Haavard Skinnemoen <hskinnemoen@gmail.com> Cc: Hans-Christian Egtvedt <egtvedt@samfundet.no> Cc: Heiko Carstens <heiko.carstens@de.ibm.com> Cc: Helge Deller <deller@gmx.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: Ivan Kokshaysky <ink@jurassic.park.msu.ru> Cc: James Hogan <james.hogan@imgtec.com> Cc: Jeff Dike <jdike@addtoit.com> Cc: Jesper Nilsson <jesper.nilsson@axis.com> Cc: Jiri Slaby <jslaby@suse.cz> Cc: Jonas Bonn <jonas@southpole.se> Cc: Koichi Yasutake <yasutake.koichi@jp.panasonic.com> Cc: Lennox Wu <lennox.wu@gmail.com> Cc: Ley Foon Tan <lftan@altera.com> Cc: Mark Salter <msalter@redhat.com> Cc: Martin Schwidefsky <schwidefsky@de.ibm.com> Cc: Matt Turner <mattst88@gmail.com> Cc: Max Filippov <jcmvbkbc@gmail.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Michal Simek <monstr@monstr.eu> Cc: Mikael Starvik <starvik@axis.com> Cc: Paul Mackerras <paulus@samba.org> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Ralf Baechle <ralf@linux-mips.org> Cc: Rich Felker <dalias@libc.org> Cc: Richard Henderson <rth@twiddle.net> Cc: Richard Kuo <rkuo@codeaurora.org> Cc: Richard Weinberger <richard@nod.at> Cc: Russell King <linux@arm.linux.org.uk> Cc: Steven Miao <realmz6@gmail.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Tony Luck <tony.luck@intel.com> Cc: Vineet Gupta <vgupta@synopsys.com> Cc: Will Deacon <will.deacon@arm.com> Cc: Yoshinori Sato <ysato@users.sourceforge.jp> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-05-21 00:00:25 +00:00
bad_fork_cleanup_thread:
exit_thread(p);
bad_fork_cleanup_io:
if (p->io_context)
exit_io_context(p);
bad_fork_cleanup_namespaces:
exit_task_namespaces(p);
bad_fork_cleanup_mm:
if (p->mm)
mmput(p->mm);
bad_fork_cleanup_signal:
clone(): fix race between copy_process() and de_thread() Spotted by Hiroshi Shimamoto who also provided the test-case below. copy_process() uses signal->count as a reference counter, but it is not. This test case #include <sys/types.h> #include <sys/wait.h> #include <unistd.h> #include <stdio.h> #include <errno.h> #include <pthread.h> void *null_thread(void *p) { for (;;) sleep(1); return NULL; } void *exec_thread(void *p) { execl("/bin/true", "/bin/true", NULL); return null_thread(p); } int main(int argc, char **argv) { for (;;) { pid_t pid; int ret, status; pid = fork(); if (pid < 0) break; if (!pid) { pthread_t tid; pthread_create(&tid, NULL, exec_thread, NULL); for (;;) pthread_create(&tid, NULL, null_thread, NULL); } do { ret = waitpid(pid, &status, 0); } while (ret == -1 && errno == EINTR); } return 0; } quickly creates an unkillable task. If copy_process(CLONE_THREAD) races with de_thread() copy_signal()->atomic(signal->count) breaks the signal->notify_count logic, and the execing thread can hang forever in kernel space. Change copy_process() to increment count/live only when we know for sure we can't fail. In this case the forked thread will take care of its reference to signal correctly. If copy_process() fails, check CLONE_THREAD flag. If it it set - do nothing, the counters were not changed and current belongs to the same thread group. If it is not set, ->signal must be released in any case (and ->count must be == 1), the forked child is the only thread in the thread group. We need more cleanups here, in particular signal->count should not be used by de_thread/__exit_signal at all. This patch only fixes the bug. Reported-by: Hiroshi Shimamoto <h-shimamoto@ct.jp.nec.com> Tested-by: Hiroshi Shimamoto <h-shimamoto@ct.jp.nec.com> Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: <stable@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-08-26 21:29:24 +00:00
if (!(clone_flags & CLONE_THREAD))
free_signal_struct(p->signal);
bad_fork_cleanup_sighand:
__cleanup_sighand(p->sighand);
bad_fork_cleanup_fs:
exit_fs(p); /* blocking */
bad_fork_cleanup_files:
exit_files(p); /* blocking */
bad_fork_cleanup_semundo:
exit_sem(p);
bad_fork_cleanup_audit:
audit_free(p);
bad_fork_cleanup_perf:
perf: Do the big rename: Performance Counters -> Performance Events Bye-bye Performance Counters, welcome Performance Events! In the past few months the perfcounters subsystem has grown out its initial role of counting hardware events, and has become (and is becoming) a much broader generic event enumeration, reporting, logging, monitoring, analysis facility. Naming its core object 'perf_counter' and naming the subsystem 'perfcounters' has become more and more of a misnomer. With pending code like hw-breakpoints support the 'counter' name is less and less appropriate. All in one, we've decided to rename the subsystem to 'performance events' and to propagate this rename through all fields, variables and API names. (in an ABI compatible fashion) The word 'event' is also a bit shorter than 'counter' - which makes it slightly more convenient to write/handle as well. Thanks goes to Stephane Eranian who first observed this misnomer and suggested a rename. User-space tooling and ABI compatibility is not affected - this patch should be function-invariant. (Also, defconfigs were not touched to keep the size down.) This patch has been generated via the following script: FILES=$(find * -type f | grep -vE 'oprofile|[^K]config') sed -i \ -e 's/PERF_EVENT_/PERF_RECORD_/g' \ -e 's/PERF_COUNTER/PERF_EVENT/g' \ -e 's/perf_counter/perf_event/g' \ -e 's/nb_counters/nb_events/g' \ -e 's/swcounter/swevent/g' \ -e 's/tpcounter_event/tp_event/g' \ $FILES for N in $(find . -name perf_counter.[ch]); do M=$(echo $N | sed 's/perf_counter/perf_event/g') mv $N $M done FILES=$(find . -name perf_event.*) sed -i \ -e 's/COUNTER_MASK/REG_MASK/g' \ -e 's/COUNTER/EVENT/g' \ -e 's/\<event\>/event_id/g' \ -e 's/counter/event/g' \ -e 's/Counter/Event/g' \ $FILES ... to keep it as correct as possible. This script can also be used by anyone who has pending perfcounters patches - it converts a Linux kernel tree over to the new naming. We tried to time this change to the point in time where the amount of pending patches is the smallest: the end of the merge window. Namespace clashes were fixed up in a preparatory patch - and some stylistic fallout will be fixed up in a subsequent patch. ( NOTE: 'counters' are still the proper terminology when we deal with hardware registers - and these sed scripts are a bit over-eager in renaming them. I've undone some of that, but in case there's something left where 'counter' would be better than 'event' we can undo that on an individual basis instead of touching an otherwise nicely automated patch. ) Suggested-by: Stephane Eranian <eranian@google.com> Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl> Acked-by: Paul Mackerras <paulus@samba.org> Reviewed-by: Arjan van de Ven <arjan@linux.intel.com> Cc: Mike Galbraith <efault@gmx.de> Cc: Arnaldo Carvalho de Melo <acme@redhat.com> Cc: Frederic Weisbecker <fweisbec@gmail.com> Cc: Steven Rostedt <rostedt@goodmis.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: David Howells <dhowells@redhat.com> Cc: Kyle McMartin <kyle@mcmartin.ca> Cc: Martin Schwidefsky <schwidefsky@de.ibm.com> Cc: "David S. Miller" <davem@davemloft.net> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: <linux-arch@vger.kernel.org> LKML-Reference: <new-submission> Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-09-21 10:02:48 +00:00
perf_event_free_task(p);
bad_fork_cleanup_policy:
#ifdef CONFIG_NUMA
mpol_put(p->mempolicy);
bad_fork_cleanup_threadgroup_lock:
#endif
delayacct_tsk_free(p);
bad_fork_cleanup_count:
CRED: Inaugurate COW credentials Inaugurate copy-on-write credentials management. This uses RCU to manage the credentials pointer in the task_struct with respect to accesses by other tasks. A process may only modify its own credentials, and so does not need locking to access or modify its own credentials. A mutex (cred_replace_mutex) is added to the task_struct to control the effect of PTRACE_ATTACHED on credential calculations, particularly with respect to execve(). With this patch, the contents of an active credentials struct may not be changed directly; rather a new set of credentials must be prepared, modified and committed using something like the following sequence of events: struct cred *new = prepare_creds(); int ret = blah(new); if (ret < 0) { abort_creds(new); return ret; } return commit_creds(new); There are some exceptions to this rule: the keyrings pointed to by the active credentials may be instantiated - keyrings violate the COW rule as managing COW keyrings is tricky, given that it is possible for a task to directly alter the keys in a keyring in use by another task. To help enforce this, various pointers to sets of credentials, such as those in the task_struct, are declared const. The purpose of this is compile-time discouragement of altering credentials through those pointers. Once a set of credentials has been made public through one of these pointers, it may not be modified, except under special circumstances: (1) Its reference count may incremented and decremented. (2) The keyrings to which it points may be modified, but not replaced. The only safe way to modify anything else is to create a replacement and commit using the functions described in Documentation/credentials.txt (which will be added by a later patch). This patch and the preceding patches have been tested with the LTP SELinux testsuite. This patch makes several logical sets of alteration: (1) execve(). This now prepares and commits credentials in various places in the security code rather than altering the current creds directly. (2) Temporary credential overrides. do_coredump() and sys_faccessat() now prepare their own credentials and temporarily override the ones currently on the acting thread, whilst preventing interference from other threads by holding cred_replace_mutex on the thread being dumped. This will be replaced in a future patch by something that hands down the credentials directly to the functions being called, rather than altering the task's objective credentials. (3) LSM interface. A number of functions have been changed, added or removed: (*) security_capset_check(), ->capset_check() (*) security_capset_set(), ->capset_set() Removed in favour of security_capset(). (*) security_capset(), ->capset() New. This is passed a pointer to the new creds, a pointer to the old creds and the proposed capability sets. It should fill in the new creds or return an error. All pointers, barring the pointer to the new creds, are now const. (*) security_bprm_apply_creds(), ->bprm_apply_creds() Changed; now returns a value, which will cause the process to be killed if it's an error. (*) security_task_alloc(), ->task_alloc_security() Removed in favour of security_prepare_creds(). (*) security_cred_free(), ->cred_free() New. Free security data attached to cred->security. (*) security_prepare_creds(), ->cred_prepare() New. Duplicate any security data attached to cred->security. (*) security_commit_creds(), ->cred_commit() New. Apply any security effects for the upcoming installation of new security by commit_creds(). (*) security_task_post_setuid(), ->task_post_setuid() Removed in favour of security_task_fix_setuid(). (*) security_task_fix_setuid(), ->task_fix_setuid() Fix up the proposed new credentials for setuid(). This is used by cap_set_fix_setuid() to implicitly adjust capabilities in line with setuid() changes. Changes are made to the new credentials, rather than the task itself as in security_task_post_setuid(). (*) security_task_reparent_to_init(), ->task_reparent_to_init() Removed. Instead the task being reparented to init is referred directly to init's credentials. NOTE! This results in the loss of some state: SELinux's osid no longer records the sid of the thread that forked it. (*) security_key_alloc(), ->key_alloc() (*) security_key_permission(), ->key_permission() Changed. These now take cred pointers rather than task pointers to refer to the security context. (4) sys_capset(). This has been simplified and uses less locking. The LSM functions it calls have been merged. (5) reparent_to_kthreadd(). This gives the current thread the same credentials as init by simply using commit_thread() to point that way. (6) __sigqueue_alloc() and switch_uid() __sigqueue_alloc() can't stop the target task from changing its creds beneath it, so this function gets a reference to the currently applicable user_struct which it then passes into the sigqueue struct it returns if successful. switch_uid() is now called from commit_creds(), and possibly should be folded into that. commit_creds() should take care of protecting __sigqueue_alloc(). (7) [sg]et[ug]id() and co and [sg]et_current_groups. The set functions now all use prepare_creds(), commit_creds() and abort_creds() to build and check a new set of credentials before applying it. security_task_set[ug]id() is called inside the prepared section. This guarantees that nothing else will affect the creds until we've finished. The calling of set_dumpable() has been moved into commit_creds(). Much of the functionality of set_user() has been moved into commit_creds(). The get functions all simply access the data directly. (8) security_task_prctl() and cap_task_prctl(). security_task_prctl() has been modified to return -ENOSYS if it doesn't want to handle a function, or otherwise return the return value directly rather than through an argument. Additionally, cap_task_prctl() now prepares a new set of credentials, even if it doesn't end up using it. (9) Keyrings. A number of changes have been made to the keyrings code: (a) switch_uid_keyring(), copy_keys(), exit_keys() and suid_keys() have all been dropped and built in to the credentials functions directly. They may want separating out again later. (b) key_alloc() and search_process_keyrings() now take a cred pointer rather than a task pointer to specify the security context. (c) copy_creds() gives a new thread within the same thread group a new thread keyring if its parent had one, otherwise it discards the thread keyring. (d) The authorisation key now points directly to the credentials to extend the search into rather pointing to the task that carries them. (e) Installing thread, process or session keyrings causes a new set of credentials to be created, even though it's not strictly necessary for process or session keyrings (they're shared). (10) Usermode helper. The usermode helper code now carries a cred struct pointer in its subprocess_info struct instead of a new session keyring pointer. This set of credentials is derived from init_cred and installed on the new process after it has been cloned. call_usermodehelper_setup() allocates the new credentials and call_usermodehelper_freeinfo() discards them if they haven't been used. A special cred function (prepare_usermodeinfo_creds()) is provided specifically for call_usermodehelper_setup() to call. call_usermodehelper_setkeys() adjusts the credentials to sport the supplied keyring as the new session keyring. (11) SELinux. SELinux has a number of changes, in addition to those to support the LSM interface changes mentioned above: (a) selinux_setprocattr() no longer does its check for whether the current ptracer can access processes with the new SID inside the lock that covers getting the ptracer's SID. Whilst this lock ensures that the check is done with the ptracer pinned, the result is only valid until the lock is released, so there's no point doing it inside the lock. (12) is_single_threaded(). This function has been extracted from selinux_setprocattr() and put into a file of its own in the lib/ directory as join_session_keyring() now wants to use it too. The code in SELinux just checked to see whether a task shared mm_structs with other tasks (CLONE_VM), but that isn't good enough. We really want to know if they're part of the same thread group (CLONE_THREAD). (13) nfsd. The NFS server daemon now has to use the COW credentials to set the credentials it is going to use. It really needs to pass the credentials down to the functions it calls, but it can't do that until other patches in this series have been applied. Signed-off-by: David Howells <dhowells@redhat.com> Acked-by: James Morris <jmorris@namei.org> Signed-off-by: James Morris <jmorris@namei.org>
2008-11-13 23:39:23 +00:00
atomic_dec(&p->cred->user->processes);
exit_creds(p);
bad_fork_free:
p->state = TASK_DEAD;
put_task_stack(p);
free_task(p);
fork_out:
return ERR_PTR(retval);
}
static inline void init_idle_pids(struct pid_link *links)
{
enum pid_type type;
for (type = PIDTYPE_PID; type < PIDTYPE_MAX; ++type) {
INIT_HLIST_NODE(&links[type].node); /* not really needed */
links[type].pid = &init_struct_pid;
}
}
struct task_struct *fork_idle(int cpu)
{
struct task_struct *task;
task = copy_process(CLONE_VM, 0, 0, NULL, &init_struct_pid, 0, 0,
cpu_to_node(cpu));
if (!IS_ERR(task)) {
init_idle_pids(task->pids);
init_idle(task, cpu);
}
return task;
}
/*
* Ok, this is the main fork-routine.
*
* It copies the process, and if successful kick-starts
* it and waits for it to finish using the VM if required.
*/
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
long _do_fork(unsigned long clone_flags,
unsigned long stack_start,
unsigned long stack_size,
int __user *parent_tidptr,
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
int __user *child_tidptr,
unsigned long tls)
{
struct task_struct *p;
int trace = 0;
[PATCH] pidhash: Refactor the pid hash table Simplifies the code, reduces the need for 4 pid hash tables, and makes the code more capable. In the discussions I had with Oleg it was felt that to a large extent the cleanup itself justified the work. With struct pid being dynamically allocated meant we could create the hash table entry when the pid was allocated and free the hash table entry when the pid was freed. Instead of playing with the hash lists when ever a process would attach or detach to a process. For myself the fact that it gave what my previous task_ref patch gave for free with simpler code was a big win. The problem is that if you hold a reference to struct task_struct you lock in 10K of low memory. If you do that in a user controllable way like /proc does, with an unprivileged but hostile user space application with typical resource limits of 1000 fds and 100 processes I can trigger the OOM killer by consuming all of low memory with task structs, on a machine wight 1GB of low memory. If I instead hold a reference to struct pid which holds a pointer to my task_struct, I don't suffer from that problem because struct pid is 2 orders of magnitude smaller. In fact struct pid is small enough that most other kernel data structures dwarf it, so simply limiting the number of referring data structures is enough to prevent exhaustion of low memory. This splits the current struct pid into two structures, struct pid and struct pid_link, and reduces our number of hash tables from PIDTYPE_MAX to just one. struct pid_link is the per process linkage into the hash tables and lives in struct task_struct. struct pid is given an indepedent lifetime, and holds pointers to each of the pid types. The independent life of struct pid simplifies attach_pid, and detach_pid, because we are always manipulating the list of pids and not the hash table. In addition in giving struct pid an indpendent life it makes the concept much more powerful. Kernel data structures can now embed a struct pid * instead of a pid_t and not suffer from pid wrap around problems or from keeping unnecessarily large amounts of memory allocated. Signed-off-by: Eric W. Biederman <ebiederm@xmission.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-03-31 10:31:42 +00:00
long nr;
/*
* Determine whether and which event to report to ptracer. When
* called from kernel_thread or CLONE_UNTRACED is explicitly
* requested, no event is reported; otherwise, report if the event
* for the type of forking is enabled.
*/
if (!(clone_flags & CLONE_UNTRACED)) {
if (clone_flags & CLONE_VFORK)
trace = PTRACE_EVENT_VFORK;
else if ((clone_flags & CSIGNAL) != SIGCHLD)
trace = PTRACE_EVENT_CLONE;
else
trace = PTRACE_EVENT_FORK;
if (likely(!ptrace_event_enabled(current, trace)))
trace = 0;
}
p = copy_process(clone_flags, stack_start, stack_size,
child_tidptr, NULL, trace, tls, NUMA_NO_NODE);
gcc-plugins: Add latent_entropy plugin This adds a new gcc plugin named "latent_entropy". It is designed to extract as much possible uncertainty from a running system at boot time as possible, hoping to capitalize on any possible variation in CPU operation (due to runtime data differences, hardware differences, SMP ordering, thermal timing variation, cache behavior, etc). At the very least, this plugin is a much more comprehensive example for how to manipulate kernel code using the gcc plugin internals. The need for very-early boot entropy tends to be very architecture or system design specific, so this plugin is more suited for those sorts of special cases. The existing kernel RNG already attempts to extract entropy from reliable runtime variation, but this plugin takes the idea to a logical extreme by permuting a global variable based on any variation in code execution (e.g. a different value (and permutation function) is used to permute the global based on loop count, case statement, if/then/else branching, etc). To do this, the plugin starts by inserting a local variable in every marked function. The plugin then adds logic so that the value of this variable is modified by randomly chosen operations (add, xor and rol) and random values (gcc generates separate static values for each location at compile time and also injects the stack pointer at runtime). The resulting value depends on the control flow path (e.g., loops and branches taken). Before the function returns, the plugin mixes this local variable into the latent_entropy global variable. The value of this global variable is added to the kernel entropy pool in do_one_initcall() and _do_fork(), though it does not credit any bytes of entropy to the pool; the contents of the global are just used to mix the pool. Additionally, the plugin can pre-initialize arrays with build-time random contents, so that two different kernel builds running on identical hardware will not have the same starting values. Signed-off-by: Emese Revfy <re.emese@gmail.com> [kees: expanded commit message and code comments] Signed-off-by: Kees Cook <keescook@chromium.org>
2016-06-20 18:41:19 +00:00
add_latent_entropy();
/*
* Do this prior waking up the new thread - the thread pointer
* might get invalid after that point, if the thread exits quickly.
*/
if (!IS_ERR(p)) {
struct completion vfork;
struct pid *pid;
trace_sched_process_fork(current, p);
pid = get_task_pid(p, PIDTYPE_PID);
nr = pid_vnr(pid);
if (clone_flags & CLONE_PARENT_SETTID)
put_user(nr, parent_tidptr);
if (clone_flags & CLONE_VFORK) {
p->vfork_done = &vfork;
init_completion(&vfork);
get_task_struct(p);
}
wake_up_new_task(p);
/* forking complete and child started to run, tell ptracer */
if (unlikely(trace))
ptrace_event_pid(trace, pid);
if (clone_flags & CLONE_VFORK) {
if (!wait_for_vfork_done(p, &vfork))
ptrace_event_pid(PTRACE_EVENT_VFORK_DONE, pid);
}
put_pid(pid);
} else {
[PATCH] pidhash: Refactor the pid hash table Simplifies the code, reduces the need for 4 pid hash tables, and makes the code more capable. In the discussions I had with Oleg it was felt that to a large extent the cleanup itself justified the work. With struct pid being dynamically allocated meant we could create the hash table entry when the pid was allocated and free the hash table entry when the pid was freed. Instead of playing with the hash lists when ever a process would attach or detach to a process. For myself the fact that it gave what my previous task_ref patch gave for free with simpler code was a big win. The problem is that if you hold a reference to struct task_struct you lock in 10K of low memory. If you do that in a user controllable way like /proc does, with an unprivileged but hostile user space application with typical resource limits of 1000 fds and 100 processes I can trigger the OOM killer by consuming all of low memory with task structs, on a machine wight 1GB of low memory. If I instead hold a reference to struct pid which holds a pointer to my task_struct, I don't suffer from that problem because struct pid is 2 orders of magnitude smaller. In fact struct pid is small enough that most other kernel data structures dwarf it, so simply limiting the number of referring data structures is enough to prevent exhaustion of low memory. This splits the current struct pid into two structures, struct pid and struct pid_link, and reduces our number of hash tables from PIDTYPE_MAX to just one. struct pid_link is the per process linkage into the hash tables and lives in struct task_struct. struct pid is given an indepedent lifetime, and holds pointers to each of the pid types. The independent life of struct pid simplifies attach_pid, and detach_pid, because we are always manipulating the list of pids and not the hash table. In addition in giving struct pid an indpendent life it makes the concept much more powerful. Kernel data structures can now embed a struct pid * instead of a pid_t and not suffer from pid wrap around problems or from keeping unnecessarily large amounts of memory allocated. Signed-off-by: Eric W. Biederman <ebiederm@xmission.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-03-31 10:31:42 +00:00
nr = PTR_ERR(p);
}
[PATCH] pidhash: Refactor the pid hash table Simplifies the code, reduces the need for 4 pid hash tables, and makes the code more capable. In the discussions I had with Oleg it was felt that to a large extent the cleanup itself justified the work. With struct pid being dynamically allocated meant we could create the hash table entry when the pid was allocated and free the hash table entry when the pid was freed. Instead of playing with the hash lists when ever a process would attach or detach to a process. For myself the fact that it gave what my previous task_ref patch gave for free with simpler code was a big win. The problem is that if you hold a reference to struct task_struct you lock in 10K of low memory. If you do that in a user controllable way like /proc does, with an unprivileged but hostile user space application with typical resource limits of 1000 fds and 100 processes I can trigger the OOM killer by consuming all of low memory with task structs, on a machine wight 1GB of low memory. If I instead hold a reference to struct pid which holds a pointer to my task_struct, I don't suffer from that problem because struct pid is 2 orders of magnitude smaller. In fact struct pid is small enough that most other kernel data structures dwarf it, so simply limiting the number of referring data structures is enough to prevent exhaustion of low memory. This splits the current struct pid into two structures, struct pid and struct pid_link, and reduces our number of hash tables from PIDTYPE_MAX to just one. struct pid_link is the per process linkage into the hash tables and lives in struct task_struct. struct pid is given an indepedent lifetime, and holds pointers to each of the pid types. The independent life of struct pid simplifies attach_pid, and detach_pid, because we are always manipulating the list of pids and not the hash table. In addition in giving struct pid an indpendent life it makes the concept much more powerful. Kernel data structures can now embed a struct pid * instead of a pid_t and not suffer from pid wrap around problems or from keeping unnecessarily large amounts of memory allocated. Signed-off-by: Eric W. Biederman <ebiederm@xmission.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-03-31 10:31:42 +00:00
return nr;
}
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
#ifndef CONFIG_HAVE_COPY_THREAD_TLS
/* For compatibility with architectures that call do_fork directly rather than
* using the syscall entry points below. */
long do_fork(unsigned long clone_flags,
unsigned long stack_start,
unsigned long stack_size,
int __user *parent_tidptr,
int __user *child_tidptr)
{
return _do_fork(clone_flags, stack_start, stack_size,
parent_tidptr, child_tidptr, 0);
}
#endif
/*
* Create a kernel thread.
*/
pid_t kernel_thread(int (*fn)(void *), void *arg, unsigned long flags)
{
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
return _do_fork(flags|CLONE_VM|CLONE_UNTRACED, (unsigned long)fn,
(unsigned long)arg, NULL, NULL, 0);
}
#ifdef __ARCH_WANT_SYS_FORK
SYSCALL_DEFINE0(fork)
{
#ifdef CONFIG_MMU
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
return _do_fork(SIGCHLD, 0, 0, NULL, NULL, 0);
#else
/* can not support in nommu mode */
return -EINVAL;
#endif
}
#endif
#ifdef __ARCH_WANT_SYS_VFORK
SYSCALL_DEFINE0(vfork)
{
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
return _do_fork(CLONE_VFORK | CLONE_VM | SIGCHLD, 0,
0, NULL, NULL, 0);
}
#endif
#ifdef __ARCH_WANT_SYS_CLONE
#ifdef CONFIG_CLONE_BACKWARDS
SYSCALL_DEFINE5(clone, unsigned long, clone_flags, unsigned long, newsp,
int __user *, parent_tidptr,
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
unsigned long, tls,
int __user *, child_tidptr)
#elif defined(CONFIG_CLONE_BACKWARDS2)
SYSCALL_DEFINE5(clone, unsigned long, newsp, unsigned long, clone_flags,
int __user *, parent_tidptr,
int __user *, child_tidptr,
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
unsigned long, tls)
#elif defined(CONFIG_CLONE_BACKWARDS3)
SYSCALL_DEFINE6(clone, unsigned long, clone_flags, unsigned long, newsp,
int, stack_size,
int __user *, parent_tidptr,
int __user *, child_tidptr,
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
unsigned long, tls)
#else
SYSCALL_DEFINE5(clone, unsigned long, clone_flags, unsigned long, newsp,
int __user *, parent_tidptr,
int __user *, child_tidptr,
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
unsigned long, tls)
#endif
{
clone: support passing tls argument via C rather than pt_regs magic clone has some of the quirkiest syscall handling in the kernel, with a pile of special cases, historical curiosities, and architecture-specific calling conventions. In particular, clone with CLONE_SETTLS accepts a parameter "tls" that the C entry point completely ignores and some assembly entry points overwrite; instead, the low-level arch-specific code pulls the tls parameter out of the arch-specific register captured as part of pt_regs on entry to the kernel. That's a massive hack, and it makes the arch-specific code only work when called via the specific existing syscall entry points; because of this hack, any new clone-like system call would have to accept an identical tls argument in exactly the same arch-specific position, rather than providing a unified system call entry point across architectures. The first patch allows architectures to handle the tls argument via normal C parameter passing, if they opt in by selecting HAVE_COPY_THREAD_TLS. The second patch makes 32-bit and 64-bit x86 opt into this. These two patches came out of the clone4 series, which isn't ready for this merge window, but these first two cleanup patches were entirely uncontroversial and have acks. I'd like to go ahead and submit these two so that other architectures can begin building on top of this and opting into HAVE_COPY_THREAD_TLS. However, I'm also happy to wait and send these through the next merge window (along with v3 of clone4) if anyone would prefer that. This patch (of 2): clone with CLONE_SETTLS accepts an argument to set the thread-local storage area for the new thread. sys_clone declares an int argument tls_val in the appropriate point in the argument list (based on the various CLONE_BACKWARDS variants), but doesn't actually use or pass along that argument. Instead, sys_clone calls do_fork, which calls copy_process, which calls the arch-specific copy_thread, and copy_thread pulls the corresponding syscall argument out of the pt_regs captured at kernel entry (knowing what argument of clone that architecture passes tls in). Apart from being awful and inscrutable, that also only works because only one code path into copy_thread can pass the CLONE_SETTLS flag, and that code path comes from sys_clone with its architecture-specific argument-passing order. This prevents introducing a new version of the clone system call without propagating the same architecture-specific position of the tls argument. However, there's no reason to pull the argument out of pt_regs when sys_clone could just pass it down via C function call arguments. Introduce a new CONFIG_HAVE_COPY_THREAD_TLS for architectures to opt into, and a new copy_thread_tls that accepts the tls parameter as an additional unsigned long (syscall-argument-sized) argument. Change sys_clone's tls argument to an unsigned long (which does not change the ABI), and pass that down to copy_thread_tls. Architectures that don't opt into copy_thread_tls will continue to ignore the C argument to sys_clone in favor of the pt_regs captured at kernel entry, and thus will be unable to introduce new versions of the clone syscall. Patch co-authored by Josh Triplett and Thiago Macieira. Signed-off-by: Josh Triplett <josh@joshtriplett.org> Acked-by: Andy Lutomirski <luto@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Thiago Macieira <thiago.macieira@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-25 22:01:19 +00:00
return _do_fork(clone_flags, newsp, 0, parent_tidptr, child_tidptr, tls);
}
#endif
void walk_process_tree(struct task_struct *top, proc_visitor visitor, void *data)
{
struct task_struct *leader, *parent, *child;
int res;
read_lock(&tasklist_lock);
leader = top = top->group_leader;
down:
for_each_thread(leader, parent) {
list_for_each_entry(child, &parent->children, sibling) {
res = visitor(child, data);
if (res) {
if (res < 0)
goto out;
leader = child;
goto down;
}
up:
;
}
}
if (leader != top) {
child = leader;
parent = child->real_parent;
leader = parent->group_leader;
goto up;
}
out:
read_unlock(&tasklist_lock);
}
#ifndef ARCH_MIN_MMSTRUCT_ALIGN
#define ARCH_MIN_MMSTRUCT_ALIGN 0
#endif
static void sighand_ctor(void *data)
{
struct sighand_struct *sighand = data;
spin_lock_init(&sighand->siglock);
init_waitqueue_head(&sighand->signalfd_wqh);
}
void __init proc_caches_init(void)
{
sighand_cachep = kmem_cache_create("sighand_cache",
sizeof(struct sighand_struct), 0,
2008-05-31 13:56:17 +00:00
SLAB_HWCACHE_ALIGN|SLAB_PANIC|SLAB_DESTROY_BY_RCU|
2016-01-14 23:18:21 +00:00
SLAB_NOTRACK|SLAB_ACCOUNT, sighand_ctor);
signal_cachep = kmem_cache_create("signal_cache",
sizeof(struct signal_struct), 0,
2016-01-14 23:18:21 +00:00
SLAB_HWCACHE_ALIGN|SLAB_PANIC|SLAB_NOTRACK|SLAB_ACCOUNT,
NULL);
files_cachep = kmem_cache_create("files_cache",
sizeof(struct files_struct), 0,
2016-01-14 23:18:21 +00:00
SLAB_HWCACHE_ALIGN|SLAB_PANIC|SLAB_NOTRACK|SLAB_ACCOUNT,
NULL);
fs_cachep = kmem_cache_create("fs_cache",
sizeof(struct fs_struct), 0,
2016-01-14 23:18:21 +00:00
SLAB_HWCACHE_ALIGN|SLAB_PANIC|SLAB_NOTRACK|SLAB_ACCOUNT,
NULL);
/*
* FIXME! The "sizeof(struct mm_struct)" currently includes the
* whole struct cpumask for the OFFSTACK case. We could change
* this to *only* allocate as much of it as required by the
* maximum number of CPU's we can ever have. The cpumask_allocation
* is at the end of the structure, exactly for that reason.
*/
mm_cachep = kmem_cache_create("mm_struct",
sizeof(struct mm_struct), ARCH_MIN_MMSTRUCT_ALIGN,
2016-01-14 23:18:21 +00:00
SLAB_HWCACHE_ALIGN|SLAB_PANIC|SLAB_NOTRACK|SLAB_ACCOUNT,
NULL);
vm_area_cachep = KMEM_CACHE(vm_area_struct, SLAB_PANIC|SLAB_ACCOUNT);
NOMMU: Make VMAs per MM as for MMU-mode linux Make VMAs per mm_struct as for MMU-mode linux. This solves two problems: (1) In SYSV SHM where nattch for a segment does not reflect the number of shmat's (and forks) done. (2) In mmap() where the VMA's vm_mm is set to point to the parent mm by an exec'ing process when VM_EXECUTABLE is specified, regardless of the fact that a VMA might be shared and already have its vm_mm assigned to another process or a dead process. A new struct (vm_region) is introduced to track a mapped region and to remember the circumstances under which it may be shared and the vm_list_struct structure is discarded as it's no longer required. This patch makes the following additional changes: (1) Regions are now allocated with alloc_pages() rather than kmalloc() and with no recourse to __GFP_COMP, so the pages are not composite. Instead, each page has a reference on it held by the region. Anything else that is interested in such a page will have to get a reference on it to retain it. When the pages are released due to unmapping, each page is passed to put_page() and will be freed when the page usage count reaches zero. (2) Excess pages are trimmed after an allocation as the allocation must be made as a power-of-2 quantity of pages. (3) VMAs are added to the parent MM's R/B tree and mmap lists. As an MM may end up with overlapping VMAs within the tree, the VMA struct address is appended to the sort key. (4) Non-anonymous VMAs are now added to the backing inode's prio list. (5) Holes may be punched in anonymous VMAs with munmap(), releasing parts of the backing region. The VMA and region structs will be split if necessary. (6) sys_shmdt() only releases one attachment to a SYSV IPC shared memory segment instead of all the attachments at that addresss. Multiple shmat()'s return the same address under NOMMU-mode instead of different virtual addresses as under MMU-mode. (7) Core dumping for ELF-FDPIC requires fewer exceptions for NOMMU-mode. (8) /proc/maps is now the global list of mapped regions, and may list bits that aren't actually mapped anywhere. (9) /proc/meminfo gains a line (tagged "MmapCopy") that indicates the amount of RAM currently allocated by mmap to hold mappable regions that can't be mapped directly. These are copies of the backing device or file if not anonymous. These changes make NOMMU mode more similar to MMU mode. The downside is that NOMMU mode requires some extra memory to track things over NOMMU without this patch (VMAs are no longer shared, and there are now region structs). Signed-off-by: David Howells <dhowells@redhat.com> Tested-by: Mike Frysinger <vapier.adi@gmail.com> Acked-by: Paul Mundt <lethal@linux-sh.org>
2009-01-08 12:04:47 +00:00
mmap_init();
nsproxy_cache_init();
}
/*
sys_unshare: remove the dead CLONE_THREAD/SIGHAND/VM code Cleanup: kill the dead code which does nothing but complicates the code and confuses the reader. sys_unshare(CLONE_THREAD/SIGHAND/VM) is not really implemented, and I doubt very much it will ever work. At least, nobody even tried since the original 99d1419d96d7df9cfa56 ("unshare system call -v5: system call handler function") was applied more than 4 years ago. And the code is not consistent. unshare_thread() always fails unconditionally, while unshare_sighand() and unshare_vm() pretend to work if there is nothing to unshare. Remove unshare_thread(), unshare_sighand(), unshare_vm() helpers and related variables and add a simple CLONE_THREAD | CLONE_SIGHAND| CLONE_VM check into check_unshare_flags(). Also, move the "CLONE_NEWNS needs CLONE_FS" check from check_unshare_flags() to sys_unshare(). This looks more consistent and matches the similar do_sysvsem check in sys_unshare(). Note: with or without this patch "atomic_read(mm->mm_users) > 1" can give a false positive due to get_task_mm(). Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: Janak Desai <janak@us.ibm.com> Cc: Daniel Lezcano <daniel.lezcano@free.fr> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-03-22 23:34:09 +00:00
* Check constraints on flags passed to the unshare system call.
*/
sys_unshare: remove the dead CLONE_THREAD/SIGHAND/VM code Cleanup: kill the dead code which does nothing but complicates the code and confuses the reader. sys_unshare(CLONE_THREAD/SIGHAND/VM) is not really implemented, and I doubt very much it will ever work. At least, nobody even tried since the original 99d1419d96d7df9cfa56 ("unshare system call -v5: system call handler function") was applied more than 4 years ago. And the code is not consistent. unshare_thread() always fails unconditionally, while unshare_sighand() and unshare_vm() pretend to work if there is nothing to unshare. Remove unshare_thread(), unshare_sighand(), unshare_vm() helpers and related variables and add a simple CLONE_THREAD | CLONE_SIGHAND| CLONE_VM check into check_unshare_flags(). Also, move the "CLONE_NEWNS needs CLONE_FS" check from check_unshare_flags() to sys_unshare(). This looks more consistent and matches the similar do_sysvsem check in sys_unshare(). Note: with or without this patch "atomic_read(mm->mm_users) > 1" can give a false positive due to get_task_mm(). Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: Janak Desai <janak@us.ibm.com> Cc: Daniel Lezcano <daniel.lezcano@free.fr> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-03-22 23:34:09 +00:00
static int check_unshare_flags(unsigned long unshare_flags)
{
sys_unshare: remove the dead CLONE_THREAD/SIGHAND/VM code Cleanup: kill the dead code which does nothing but complicates the code and confuses the reader. sys_unshare(CLONE_THREAD/SIGHAND/VM) is not really implemented, and I doubt very much it will ever work. At least, nobody even tried since the original 99d1419d96d7df9cfa56 ("unshare system call -v5: system call handler function") was applied more than 4 years ago. And the code is not consistent. unshare_thread() always fails unconditionally, while unshare_sighand() and unshare_vm() pretend to work if there is nothing to unshare. Remove unshare_thread(), unshare_sighand(), unshare_vm() helpers and related variables and add a simple CLONE_THREAD | CLONE_SIGHAND| CLONE_VM check into check_unshare_flags(). Also, move the "CLONE_NEWNS needs CLONE_FS" check from check_unshare_flags() to sys_unshare(). This looks more consistent and matches the similar do_sysvsem check in sys_unshare(). Note: with or without this patch "atomic_read(mm->mm_users) > 1" can give a false positive due to get_task_mm(). Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: Janak Desai <janak@us.ibm.com> Cc: Daniel Lezcano <daniel.lezcano@free.fr> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-03-22 23:34:09 +00:00
if (unshare_flags & ~(CLONE_THREAD|CLONE_FS|CLONE_NEWNS|CLONE_SIGHAND|
CLONE_VM|CLONE_FILES|CLONE_SYSVSEM|
pidns: Support unsharing the pid namespace. Unsharing of the pid namespace unlike unsharing of other namespaces does not take affect immediately. Instead it affects the children created with fork and clone. The first of these children becomes the init process of the new pid namespace, the rest become oddball children of pid 0. From the point of view of the new pid namespace the process that created it is pid 0, as it's pid does not map. A couple of different semantics were considered but this one was settled on because it is easy to implement and it is usable from pam modules. The core reasons for the existence of unshare. I took a survey of the callers of pam modules and the following appears to be a representative sample of their logic. { setup stuff include pam child = fork(); if (!child) { setuid() exec /bin/bash } waitpid(child); pam and other cleanup } As you can see there is a fork to create the unprivileged user space process. Which means that the unprivileged user space process will appear as pid 1 in the new pid namespace. Further most login processes do not cope with extraneous children which means shifting the duty of reaping extraneous child process to the creator of those extraneous children makes the system more comprehensible. The practical reason for this set of pid namespace semantics is that it is simple to implement and verify they work correctly. Whereas an implementation that requres changing the struct pid on a process comes with a lot more races and pain. Not the least of which is that glibc caches getpid(). These semantics are implemented by having two notions of the pid namespace of a proces. There is task_active_pid_ns which is the pid namspace the process was created with and the pid namespace that all pids are presented to that process in. The task_active_pid_ns is stored in the struct pid of the task. Then there is the pid namespace that will be used for children that pid namespace is stored in task->nsproxy->pid_ns. Signed-off-by: Eric W. Biederman <ebiederm@xmission.com>
2010-03-02 23:41:50 +00:00
CLONE_NEWUTS|CLONE_NEWIPC|CLONE_NEWNET|
CLONE_NEWUSER|CLONE_NEWPID|CLONE_NEWCGROUP))
sys_unshare: remove the dead CLONE_THREAD/SIGHAND/VM code Cleanup: kill the dead code which does nothing but complicates the code and confuses the reader. sys_unshare(CLONE_THREAD/SIGHAND/VM) is not really implemented, and I doubt very much it will ever work. At least, nobody even tried since the original 99d1419d96d7df9cfa56 ("unshare system call -v5: system call handler function") was applied more than 4 years ago. And the code is not consistent. unshare_thread() always fails unconditionally, while unshare_sighand() and unshare_vm() pretend to work if there is nothing to unshare. Remove unshare_thread(), unshare_sighand(), unshare_vm() helpers and related variables and add a simple CLONE_THREAD | CLONE_SIGHAND| CLONE_VM check into check_unshare_flags(). Also, move the "CLONE_NEWNS needs CLONE_FS" check from check_unshare_flags() to sys_unshare(). This looks more consistent and matches the similar do_sysvsem check in sys_unshare(). Note: with or without this patch "atomic_read(mm->mm_users) > 1" can give a false positive due to get_task_mm(). Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: Janak Desai <janak@us.ibm.com> Cc: Daniel Lezcano <daniel.lezcano@free.fr> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-03-22 23:34:09 +00:00
return -EINVAL;
/*
unshare: Unsharing a thread does not require unsharing a vm In the logic in the initial commit of unshare made creating a new thread group for a process, contingent upon creating a new memory address space for that process. That is wrong. Two separate processes in different thread groups can share a memory address space and clone allows creation of such proceses. This is significant because it was observed that mm_users > 1 does not mean that a process is multi-threaded, as reading /proc/PID/maps temporarily increments mm_users, which allows other processes to (accidentally) interfere with unshare() calls. Correct the check in check_unshare_flags() to test for !thread_group_empty() for CLONE_THREAD, CLONE_SIGHAND, and CLONE_VM. For sighand->count > 1 for CLONE_SIGHAND and CLONE_VM. For !current_is_single_threaded instead of mm_users > 1 for CLONE_VM. By using the correct checks in unshare this removes the possibility of an accidental denial of service attack. Additionally using the correct checks in unshare ensures that only an explicit unshare(CLONE_VM) can possibly trigger the slow path of current_is_single_threaded(). As an explict unshare(CLONE_VM) is pointless it is not expected there are many applications that make that call. Cc: stable@vger.kernel.org Fixes: b2e0d98705e60e45bbb3c0032c48824ad7ae0704 userns: Implement unshare of the user namespace Reported-by: Ricky Zhou <rickyz@chromium.org> Reported-by: Kees Cook <keescook@chromium.org> Reviewed-by: Kees Cook <keescook@chromium.org> Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2015-08-10 22:35:07 +00:00
* Not implemented, but pretend it works if there is nothing
* to unshare. Note that unsharing the address space or the
* signal handlers also need to unshare the signal queues (aka
* CLONE_THREAD).
*/
sys_unshare: remove the dead CLONE_THREAD/SIGHAND/VM code Cleanup: kill the dead code which does nothing but complicates the code and confuses the reader. sys_unshare(CLONE_THREAD/SIGHAND/VM) is not really implemented, and I doubt very much it will ever work. At least, nobody even tried since the original 99d1419d96d7df9cfa56 ("unshare system call -v5: system call handler function") was applied more than 4 years ago. And the code is not consistent. unshare_thread() always fails unconditionally, while unshare_sighand() and unshare_vm() pretend to work if there is nothing to unshare. Remove unshare_thread(), unshare_sighand(), unshare_vm() helpers and related variables and add a simple CLONE_THREAD | CLONE_SIGHAND| CLONE_VM check into check_unshare_flags(). Also, move the "CLONE_NEWNS needs CLONE_FS" check from check_unshare_flags() to sys_unshare(). This looks more consistent and matches the similar do_sysvsem check in sys_unshare(). Note: with or without this patch "atomic_read(mm->mm_users) > 1" can give a false positive due to get_task_mm(). Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: Janak Desai <janak@us.ibm.com> Cc: Daniel Lezcano <daniel.lezcano@free.fr> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-03-22 23:34:09 +00:00
if (unshare_flags & (CLONE_THREAD | CLONE_SIGHAND | CLONE_VM)) {
unshare: Unsharing a thread does not require unsharing a vm In the logic in the initial commit of unshare made creating a new thread group for a process, contingent upon creating a new memory address space for that process. That is wrong. Two separate processes in different thread groups can share a memory address space and clone allows creation of such proceses. This is significant because it was observed that mm_users > 1 does not mean that a process is multi-threaded, as reading /proc/PID/maps temporarily increments mm_users, which allows other processes to (accidentally) interfere with unshare() calls. Correct the check in check_unshare_flags() to test for !thread_group_empty() for CLONE_THREAD, CLONE_SIGHAND, and CLONE_VM. For sighand->count > 1 for CLONE_SIGHAND and CLONE_VM. For !current_is_single_threaded instead of mm_users > 1 for CLONE_VM. By using the correct checks in unshare this removes the possibility of an accidental denial of service attack. Additionally using the correct checks in unshare ensures that only an explicit unshare(CLONE_VM) can possibly trigger the slow path of current_is_single_threaded(). As an explict unshare(CLONE_VM) is pointless it is not expected there are many applications that make that call. Cc: stable@vger.kernel.org Fixes: b2e0d98705e60e45bbb3c0032c48824ad7ae0704 userns: Implement unshare of the user namespace Reported-by: Ricky Zhou <rickyz@chromium.org> Reported-by: Kees Cook <keescook@chromium.org> Reviewed-by: Kees Cook <keescook@chromium.org> Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2015-08-10 22:35:07 +00:00
if (!thread_group_empty(current))
return -EINVAL;
}
if (unshare_flags & (CLONE_SIGHAND | CLONE_VM)) {
if (atomic_read(&current->sighand->count) > 1)
return -EINVAL;
}
if (unshare_flags & CLONE_VM) {
if (!current_is_single_threaded())
sys_unshare: remove the dead CLONE_THREAD/SIGHAND/VM code Cleanup: kill the dead code which does nothing but complicates the code and confuses the reader. sys_unshare(CLONE_THREAD/SIGHAND/VM) is not really implemented, and I doubt very much it will ever work. At least, nobody even tried since the original 99d1419d96d7df9cfa56 ("unshare system call -v5: system call handler function") was applied more than 4 years ago. And the code is not consistent. unshare_thread() always fails unconditionally, while unshare_sighand() and unshare_vm() pretend to work if there is nothing to unshare. Remove unshare_thread(), unshare_sighand(), unshare_vm() helpers and related variables and add a simple CLONE_THREAD | CLONE_SIGHAND| CLONE_VM check into check_unshare_flags(). Also, move the "CLONE_NEWNS needs CLONE_FS" check from check_unshare_flags() to sys_unshare(). This looks more consistent and matches the similar do_sysvsem check in sys_unshare(). Note: with or without this patch "atomic_read(mm->mm_users) > 1" can give a false positive due to get_task_mm(). Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: Janak Desai <janak@us.ibm.com> Cc: Daniel Lezcano <daniel.lezcano@free.fr> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-03-22 23:34:09 +00:00
return -EINVAL;
}
return 0;
}
/*
* Unshare the filesystem structure if it is being shared
*/
static int unshare_fs(unsigned long unshare_flags, struct fs_struct **new_fsp)
{
struct fs_struct *fs = current->fs;
if (!(unshare_flags & CLONE_FS) || !fs)
return 0;
/* don't need lock here; in the worst case we'll do useless copy */
if (fs->users == 1)
return 0;
*new_fsp = copy_fs_struct(fs);
if (!*new_fsp)
return -ENOMEM;
return 0;
}
/*
* Unshare file descriptor table if it is being shared
*/
static int unshare_fd(unsigned long unshare_flags, struct files_struct **new_fdp)
{
struct files_struct *fd = current->files;
int error = 0;
if ((unshare_flags & CLONE_FILES) &&
(fd && atomic_read(&fd->count) > 1)) {
*new_fdp = dup_fd(fd, &error);
if (!*new_fdp)
return error;
}
return 0;
}
/*
* unshare allows a process to 'unshare' part of the process
* context which was originally shared using clone. copy_*
* functions used by do_fork() cannot be used here directly
* because they modify an inactive task_struct that is being
* constructed. Here we are modifying the current, active,
* task_struct.
*/
SYSCALL_DEFINE1(unshare, unsigned long, unshare_flags)
{
struct fs_struct *fs, *new_fs = NULL;
struct files_struct *fd, *new_fd = NULL;
struct cred *new_cred = NULL;
Make access to task's nsproxy lighter When someone wants to deal with some other taks's namespaces it has to lock the task and then to get the desired namespace if the one exists. This is slow on read-only paths and may be impossible in some cases. E.g. Oleg recently noticed a race between unshare() and the (sent for review in cgroups) pid namespaces - when the task notifies the parent it has to know the parent's namespace, but taking the task_lock() is impossible there - the code is under write locked tasklist lock. On the other hand switching the namespace on task (daemonize) and releasing the namespace (after the last task exit) is rather rare operation and we can sacrifice its speed to solve the issues above. The access to other task namespaces is proposed to be performed like this: rcu_read_lock(); nsproxy = task_nsproxy(tsk); if (nsproxy != NULL) { / * * work with the namespaces here * e.g. get the reference on one of them * / } / * * NULL task_nsproxy() means that this task is * almost dead (zombie) * / rcu_read_unlock(); This patch has passed the review by Eric and Oleg :) and, of course, tested. [clg@fr.ibm.com: fix unshare()] [ebiederm@xmission.com: Update get_net_ns_by_pid] Signed-off-by: Pavel Emelyanov <xemul@openvz.org> Signed-off-by: Eric W. Biederman <ebiederm@xmission.com> Cc: Oleg Nesterov <oleg@tv-sign.ru> Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com> Cc: Serge Hallyn <serue@us.ibm.com> Signed-off-by: Cedric Le Goater <clg@fr.ibm.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-19 06:39:54 +00:00
struct nsproxy *new_nsproxy = NULL;
int do_sysvsem = 0;
sys_unshare: remove the dead CLONE_THREAD/SIGHAND/VM code Cleanup: kill the dead code which does nothing but complicates the code and confuses the reader. sys_unshare(CLONE_THREAD/SIGHAND/VM) is not really implemented, and I doubt very much it will ever work. At least, nobody even tried since the original 99d1419d96d7df9cfa56 ("unshare system call -v5: system call handler function") was applied more than 4 years ago. And the code is not consistent. unshare_thread() always fails unconditionally, while unshare_sighand() and unshare_vm() pretend to work if there is nothing to unshare. Remove unshare_thread(), unshare_sighand(), unshare_vm() helpers and related variables and add a simple CLONE_THREAD | CLONE_SIGHAND| CLONE_VM check into check_unshare_flags(). Also, move the "CLONE_NEWNS needs CLONE_FS" check from check_unshare_flags() to sys_unshare(). This looks more consistent and matches the similar do_sysvsem check in sys_unshare(). Note: with or without this patch "atomic_read(mm->mm_users) > 1" can give a false positive due to get_task_mm(). Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: Janak Desai <janak@us.ibm.com> Cc: Daniel Lezcano <daniel.lezcano@free.fr> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-03-22 23:34:09 +00:00
int err;
/*
* If unsharing a user namespace must also unshare the thread group
* and unshare the filesystem root and working directories.
*/
if (unshare_flags & CLONE_NEWUSER)
unshare_flags |= CLONE_THREAD | CLONE_FS;
pidns: Support unsharing the pid namespace. Unsharing of the pid namespace unlike unsharing of other namespaces does not take affect immediately. Instead it affects the children created with fork and clone. The first of these children becomes the init process of the new pid namespace, the rest become oddball children of pid 0. From the point of view of the new pid namespace the process that created it is pid 0, as it's pid does not map. A couple of different semantics were considered but this one was settled on because it is easy to implement and it is usable from pam modules. The core reasons for the existence of unshare. I took a survey of the callers of pam modules and the following appears to be a representative sample of their logic. { setup stuff include pam child = fork(); if (!child) { setuid() exec /bin/bash } waitpid(child); pam and other cleanup } As you can see there is a fork to create the unprivileged user space process. Which means that the unprivileged user space process will appear as pid 1 in the new pid namespace. Further most login processes do not cope with extraneous children which means shifting the duty of reaping extraneous child process to the creator of those extraneous children makes the system more comprehensible. The practical reason for this set of pid namespace semantics is that it is simple to implement and verify they work correctly. Whereas an implementation that requres changing the struct pid on a process comes with a lot more races and pain. Not the least of which is that glibc caches getpid(). These semantics are implemented by having two notions of the pid namespace of a proces. There is task_active_pid_ns which is the pid namspace the process was created with and the pid namespace that all pids are presented to that process in. The task_active_pid_ns is stored in the struct pid of the task. Then there is the pid namespace that will be used for children that pid namespace is stored in task->nsproxy->pid_ns. Signed-off-by: Eric W. Biederman <ebiederm@xmission.com>
2010-03-02 23:41:50 +00:00
/*
* If unsharing vm, must also unshare signal handlers.
*/
if (unshare_flags & CLONE_VM)
unshare_flags |= CLONE_SIGHAND;
unshare: Unsharing a thread does not require unsharing a vm In the logic in the initial commit of unshare made creating a new thread group for a process, contingent upon creating a new memory address space for that process. That is wrong. Two separate processes in different thread groups can share a memory address space and clone allows creation of such proceses. This is significant because it was observed that mm_users > 1 does not mean that a process is multi-threaded, as reading /proc/PID/maps temporarily increments mm_users, which allows other processes to (accidentally) interfere with unshare() calls. Correct the check in check_unshare_flags() to test for !thread_group_empty() for CLONE_THREAD, CLONE_SIGHAND, and CLONE_VM. For sighand->count > 1 for CLONE_SIGHAND and CLONE_VM. For !current_is_single_threaded instead of mm_users > 1 for CLONE_VM. By using the correct checks in unshare this removes the possibility of an accidental denial of service attack. Additionally using the correct checks in unshare ensures that only an explicit unshare(CLONE_VM) can possibly trigger the slow path of current_is_single_threaded(). As an explict unshare(CLONE_VM) is pointless it is not expected there are many applications that make that call. Cc: stable@vger.kernel.org Fixes: b2e0d98705e60e45bbb3c0032c48824ad7ae0704 userns: Implement unshare of the user namespace Reported-by: Ricky Zhou <rickyz@chromium.org> Reported-by: Kees Cook <keescook@chromium.org> Reviewed-by: Kees Cook <keescook@chromium.org> Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2015-08-10 22:35:07 +00:00
/*
* If unsharing a signal handlers, must also unshare the signal queues.
*/
if (unshare_flags & CLONE_SIGHAND)
unshare_flags |= CLONE_THREAD;
sys_unshare: remove the dead CLONE_THREAD/SIGHAND/VM code Cleanup: kill the dead code which does nothing but complicates the code and confuses the reader. sys_unshare(CLONE_THREAD/SIGHAND/VM) is not really implemented, and I doubt very much it will ever work. At least, nobody even tried since the original 99d1419d96d7df9cfa56 ("unshare system call -v5: system call handler function") was applied more than 4 years ago. And the code is not consistent. unshare_thread() always fails unconditionally, while unshare_sighand() and unshare_vm() pretend to work if there is nothing to unshare. Remove unshare_thread(), unshare_sighand(), unshare_vm() helpers and related variables and add a simple CLONE_THREAD | CLONE_SIGHAND| CLONE_VM check into check_unshare_flags(). Also, move the "CLONE_NEWNS needs CLONE_FS" check from check_unshare_flags() to sys_unshare(). This looks more consistent and matches the similar do_sysvsem check in sys_unshare(). Note: with or without this patch "atomic_read(mm->mm_users) > 1" can give a false positive due to get_task_mm(). Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: Janak Desai <janak@us.ibm.com> Cc: Daniel Lezcano <daniel.lezcano@free.fr> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-03-22 23:34:09 +00:00
/*
* If unsharing namespace, must also unshare filesystem information.
*/
if (unshare_flags & CLONE_NEWNS)
unshare_flags |= CLONE_FS;
pidns: Support unsharing the pid namespace. Unsharing of the pid namespace unlike unsharing of other namespaces does not take affect immediately. Instead it affects the children created with fork and clone. The first of these children becomes the init process of the new pid namespace, the rest become oddball children of pid 0. From the point of view of the new pid namespace the process that created it is pid 0, as it's pid does not map. A couple of different semantics were considered but this one was settled on because it is easy to implement and it is usable from pam modules. The core reasons for the existence of unshare. I took a survey of the callers of pam modules and the following appears to be a representative sample of their logic. { setup stuff include pam child = fork(); if (!child) { setuid() exec /bin/bash } waitpid(child); pam and other cleanup } As you can see there is a fork to create the unprivileged user space process. Which means that the unprivileged user space process will appear as pid 1 in the new pid namespace. Further most login processes do not cope with extraneous children which means shifting the duty of reaping extraneous child process to the creator of those extraneous children makes the system more comprehensible. The practical reason for this set of pid namespace semantics is that it is simple to implement and verify they work correctly. Whereas an implementation that requres changing the struct pid on a process comes with a lot more races and pain. Not the least of which is that glibc caches getpid(). These semantics are implemented by having two notions of the pid namespace of a proces. There is task_active_pid_ns which is the pid namspace the process was created with and the pid namespace that all pids are presented to that process in. The task_active_pid_ns is stored in the struct pid of the task. Then there is the pid namespace that will be used for children that pid namespace is stored in task->nsproxy->pid_ns. Signed-off-by: Eric W. Biederman <ebiederm@xmission.com>
2010-03-02 23:41:50 +00:00
err = check_unshare_flags(unshare_flags);
if (err)
goto bad_unshare_out;
/*
* CLONE_NEWIPC must also detach from the undolist: after switching
* to a new ipc namespace, the semaphore arrays from the old
* namespace are unreachable.
*/
if (unshare_flags & (CLONE_NEWIPC|CLONE_SYSVSEM))
do_sysvsem = 1;
err = unshare_fs(unshare_flags, &new_fs);
if (err)
sys_unshare: remove the dead CLONE_THREAD/SIGHAND/VM code Cleanup: kill the dead code which does nothing but complicates the code and confuses the reader. sys_unshare(CLONE_THREAD/SIGHAND/VM) is not really implemented, and I doubt very much it will ever work. At least, nobody even tried since the original 99d1419d96d7df9cfa56 ("unshare system call -v5: system call handler function") was applied more than 4 years ago. And the code is not consistent. unshare_thread() always fails unconditionally, while unshare_sighand() and unshare_vm() pretend to work if there is nothing to unshare. Remove unshare_thread(), unshare_sighand(), unshare_vm() helpers and related variables and add a simple CLONE_THREAD | CLONE_SIGHAND| CLONE_VM check into check_unshare_flags(). Also, move the "CLONE_NEWNS needs CLONE_FS" check from check_unshare_flags() to sys_unshare(). This looks more consistent and matches the similar do_sysvsem check in sys_unshare(). Note: with or without this patch "atomic_read(mm->mm_users) > 1" can give a false positive due to get_task_mm(). Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: Janak Desai <janak@us.ibm.com> Cc: Daniel Lezcano <daniel.lezcano@free.fr> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-03-22 23:34:09 +00:00
goto bad_unshare_out;
err = unshare_fd(unshare_flags, &new_fd);
if (err)
sys_unshare: remove the dead CLONE_THREAD/SIGHAND/VM code Cleanup: kill the dead code which does nothing but complicates the code and confuses the reader. sys_unshare(CLONE_THREAD/SIGHAND/VM) is not really implemented, and I doubt very much it will ever work. At least, nobody even tried since the original 99d1419d96d7df9cfa56 ("unshare system call -v5: system call handler function") was applied more than 4 years ago. And the code is not consistent. unshare_thread() always fails unconditionally, while unshare_sighand() and unshare_vm() pretend to work if there is nothing to unshare. Remove unshare_thread(), unshare_sighand(), unshare_vm() helpers and related variables and add a simple CLONE_THREAD | CLONE_SIGHAND| CLONE_VM check into check_unshare_flags(). Also, move the "CLONE_NEWNS needs CLONE_FS" check from check_unshare_flags() to sys_unshare(). This looks more consistent and matches the similar do_sysvsem check in sys_unshare(). Note: with or without this patch "atomic_read(mm->mm_users) > 1" can give a false positive due to get_task_mm(). Signed-off-by: Oleg Nesterov <oleg@redhat.com> Acked-by: Roland McGrath <roland@redhat.com> Cc: Janak Desai <janak@us.ibm.com> Cc: Daniel Lezcano <daniel.lezcano@free.fr> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Acked-by: Serge Hallyn <serge.hallyn@canonical.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-03-22 23:34:09 +00:00
goto bad_unshare_cleanup_fs;
err = unshare_userns(unshare_flags, &new_cred);
if (err)
goto bad_unshare_cleanup_fd;
err = unshare_nsproxy_namespaces(unshare_flags, &new_nsproxy,
new_cred, new_fs);
if (err)
goto bad_unshare_cleanup_cred;
if (new_fs || new_fd || do_sysvsem || new_cred || new_nsproxy) {
if (do_sysvsem) {
/*
* CLONE_SYSVSEM is equivalent to sys_exit().
*/
exit_sem(current);
}
shm: make exit_shm work proportional to task activity This is small set of patches our team has had kicking around for a few versions internally that fixes tasks getting hung on shm_exit when there are many threads hammering it at once. Anton wrote a simple test to cause the issue: http://ozlabs.org/~anton/junkcode/bust_shm_exit.c Before applying this patchset, this test code will cause either hanging tracebacks or pthread out of memory errors. After this patchset, it will still produce output like: root@somehost:~# ./bust_shm_exit 1024 160 ... INFO: rcu_sched detected stalls on CPUs/tasks: {} (detected by 116, t=2111 jiffies, g=241, c=240, q=7113) INFO: Stall ended before state dump start ... But the task will continue to run along happily, so we consider this an improvement over hanging, even if it's a bit noisy. This patch (of 3): exit_shm obtains the ipc_ns shm rwsem for write and holds it while it walks every shared memory segment in the namespace. Thus the amount of work is related to the number of shm segments in the namespace not the number of segments that might need to be cleaned. In addition, this occurs after the task has been notified the thread has exited, so the number of tasks waiting for the ns shm rwsem can grow without bound until memory is exausted. Add a list to the task struct of all shmids allocated by this task. Init the list head in copy_process. Use the ns->rwsem for locking. Add segments after id is added, remove before removing from id. On unshare of NEW_IPCNS orphan any ids as if the task had exited, similar to handling of semaphore undo. I chose a define for the init sequence since its a simple list init, otherwise it would require a function call to avoid include loops between the semaphore code and the task struct. Converting the list_del to list_del_init for the unshare cases would remove the exit followed by init, but I left it blow up if not inited. Signed-off-by: Milton Miller <miltonm@bga.com> Signed-off-by: Jack Miller <millerjo@us.ibm.com> Cc: Davidlohr Bueso <davidlohr@hp.com> Cc: Manfred Spraul <manfred@colorfullife.com> Cc: Anton Blanchard <anton@samba.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-08 21:23:19 +00:00
if (unshare_flags & CLONE_NEWIPC) {
/* Orphan segments in old ns (see sem above). */
exit_shm(current);
shm_init_task(current);
}
if (new_nsproxy)
Make access to task's nsproxy lighter When someone wants to deal with some other taks's namespaces it has to lock the task and then to get the desired namespace if the one exists. This is slow on read-only paths and may be impossible in some cases. E.g. Oleg recently noticed a race between unshare() and the (sent for review in cgroups) pid namespaces - when the task notifies the parent it has to know the parent's namespace, but taking the task_lock() is impossible there - the code is under write locked tasklist lock. On the other hand switching the namespace on task (daemonize) and releasing the namespace (after the last task exit) is rather rare operation and we can sacrifice its speed to solve the issues above. The access to other task namespaces is proposed to be performed like this: rcu_read_lock(); nsproxy = task_nsproxy(tsk); if (nsproxy != NULL) { / * * work with the namespaces here * e.g. get the reference on one of them * / } / * * NULL task_nsproxy() means that this task is * almost dead (zombie) * / rcu_read_unlock(); This patch has passed the review by Eric and Oleg :) and, of course, tested. [clg@fr.ibm.com: fix unshare()] [ebiederm@xmission.com: Update get_net_ns_by_pid] Signed-off-by: Pavel Emelyanov <xemul@openvz.org> Signed-off-by: Eric W. Biederman <ebiederm@xmission.com> Cc: Oleg Nesterov <oleg@tv-sign.ru> Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com> Cc: Serge Hallyn <serue@us.ibm.com> Signed-off-by: Cedric Le Goater <clg@fr.ibm.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-19 06:39:54 +00:00
switch_task_namespaces(current, new_nsproxy);
Make access to task's nsproxy lighter When someone wants to deal with some other taks's namespaces it has to lock the task and then to get the desired namespace if the one exists. This is slow on read-only paths and may be impossible in some cases. E.g. Oleg recently noticed a race between unshare() and the (sent for review in cgroups) pid namespaces - when the task notifies the parent it has to know the parent's namespace, but taking the task_lock() is impossible there - the code is under write locked tasklist lock. On the other hand switching the namespace on task (daemonize) and releasing the namespace (after the last task exit) is rather rare operation and we can sacrifice its speed to solve the issues above. The access to other task namespaces is proposed to be performed like this: rcu_read_lock(); nsproxy = task_nsproxy(tsk); if (nsproxy != NULL) { / * * work with the namespaces here * e.g. get the reference on one of them * / } / * * NULL task_nsproxy() means that this task is * almost dead (zombie) * / rcu_read_unlock(); This patch has passed the review by Eric and Oleg :) and, of course, tested. [clg@fr.ibm.com: fix unshare()] [ebiederm@xmission.com: Update get_net_ns_by_pid] Signed-off-by: Pavel Emelyanov <xemul@openvz.org> Signed-off-by: Eric W. Biederman <ebiederm@xmission.com> Cc: Oleg Nesterov <oleg@tv-sign.ru> Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com> Cc: Serge Hallyn <serue@us.ibm.com> Signed-off-by: Cedric Le Goater <clg@fr.ibm.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-19 06:39:54 +00:00
task_lock(current);
if (new_fs) {
fs = current->fs;
spin_lock(&fs->lock);
current->fs = new_fs;
if (--fs->users)
new_fs = NULL;
else
new_fs = fs;
spin_unlock(&fs->lock);
}
if (new_fd) {
fd = current->files;
current->files = new_fd;
new_fd = fd;
}
task_unlock(current);
if (new_cred) {
/* Install the new user namespace */
commit_creds(new_cred);
new_cred = NULL;
}
}
bad_unshare_cleanup_cred:
if (new_cred)
put_cred(new_cred);
bad_unshare_cleanup_fd:
if (new_fd)
put_files_struct(new_fd);
bad_unshare_cleanup_fs:
if (new_fs)
free_fs_struct(new_fs);
bad_unshare_out:
return err;
}
/*
* Helper to unshare the files of the current task.
* We don't want to expose copy_files internals to
* the exec layer of the kernel.
*/
int unshare_files(struct files_struct **displaced)
{
struct task_struct *task = current;
struct files_struct *copy = NULL;
int error;
error = unshare_fd(CLONE_FILES, &copy);
if (error || !copy) {
*displaced = NULL;
return error;
}
*displaced = task->files;
task_lock(task);
task->files = copy;
task_unlock(task);
return 0;
}
int sysctl_max_threads(struct ctl_table *table, int write,
void __user *buffer, size_t *lenp, loff_t *ppos)
{
struct ctl_table t;
int ret;
int threads = max_threads;
int min = MIN_THREADS;
int max = MAX_THREADS;
t = *table;
t.data = &threads;
t.extra1 = &min;
t.extra2 = &max;
ret = proc_dointvec_minmax(&t, write, buffer, lenp, ppos);
if (ret || !write)
return ret;
set_max_threads(threads);
return 0;
}