linux/fs/file_table.c

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// SPDX-License-Identifier: GPL-2.0-only
/*
* linux/fs/file_table.c
*
* Copyright (C) 1991, 1992 Linus Torvalds
* Copyright (C) 1997 David S. Miller (davem@caip.rutgers.edu)
*/
#include <linux/string.h>
#include <linux/slab.h>
#include <linux/file.h>
#include <linux/init.h>
#include <linux/module.h>
#include <linux/fs.h>
#include <linux/filelock.h>
#include <linux/security.h>
#include <linux/cred.h>
#include <linux/eventpoll.h>
#include <linux/rcupdate.h>
#include <linux/mount.h>
#include <linux/capability.h>
#include <linux/cdev.h>
#include <linux/fsnotify.h>
#include <linux/sysctl.h>
#include <linux/percpu_counter.h>
fs: scale files_lock fs: scale files_lock Improve scalability of files_lock by adding per-cpu, per-sb files lists, protected with an lglock. The lglock provides fast access to the per-cpu lists to add and remove files. It also provides a snapshot of all the per-cpu lists (although this is very slow). One difficulty with this approach is that a file can be removed from the list by another CPU. We must track which per-cpu list the file is on with a new variale in the file struct (packed into a hole on 64-bit archs). Scalability could suffer if files are frequently removed from different cpu's list. However loads with frequent removal of files imply short interval between adding and removing the files, and the scheduler attempts to avoid moving processes too far away. Also, even in the case of cross-CPU removal, the hardware has much more opportunity to parallelise cacheline transfers with N cachelines than with 1. A worst-case test of 1 CPU allocating files subsequently being freed by N CPUs degenerates to contending on a single lock, which is no worse than before. When more than one CPU are allocating files, even if they are always freed by different CPUs, there will be more parallelism than the single-lock case. Testing results: On a 2 socket, 8 core opteron, I measure the number of times the lock is taken to remove the file, the number of times it is removed by the same CPU that added it, and the number of times it is removed by the same node that added it. Booting: locks= 25049 cpu-hits= 23174 (92.5%) node-hits= 23945 (95.6%) kbuild -j16 locks=2281913 cpu-hits=2208126 (96.8%) node-hits=2252674 (98.7%) dbench 64 locks=4306582 cpu-hits=4287247 (99.6%) node-hits=4299527 (99.8%) So a file is removed from the same CPU it was added by over 90% of the time. It remains within the same node 95% of the time. Tim Chen ran some numbers for a 64 thread Nehalem system performing a compile. throughput 2.6.34-rc2 24.5 +patch 24.9 us sys idle IO wait (in %) 2.6.34-rc2 51.25 28.25 17.25 3.25 +patch 53.75 18.5 19 8.75 So significantly less CPU time spent in kernel code, higher idle time and slightly higher throughput. Single threaded performance difference was within the noise of microbenchmarks. That is not to say penalty does not exist, the code is larger and more memory accesses required so it will be slightly slower. Cc: linux-kernel@vger.kernel.org Cc: Tim Chen <tim.c.chen@linux.intel.com> Cc: Andi Kleen <ak@linux.intel.com> Signed-off-by: Nick Piggin <npiggin@kernel.dk> Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-17 18:37:38 +00:00
#include <linux/percpu.h>
#include <linux/task_work.h>
#include <linux/swap.h>
#include <linux/kmemleak.h>
#include <linux/atomic.h>
#include "internal.h"
/* sysctl tunables... */
static struct files_stat_struct files_stat = {
.max_files = NR_FILE
};
/* SLAB cache for file structures */
static struct kmem_cache *filp_cachep __ro_after_init;
static struct kmem_cache *bfilp_cachep __ro_after_init;
static struct percpu_counter nr_files __cacheline_aligned_in_smp;
/* Container for backing file with optional user path */
struct backing_file {
struct file file;
union {
struct path user_path;
freeptr_t bf_freeptr;
};
};
static inline struct backing_file *backing_file(struct file *f)
{
return container_of(f, struct backing_file, file);
}
struct path *backing_file_user_path(struct file *f)
{
return &backing_file(f)->user_path;
}
EXPORT_SYMBOL_GPL(backing_file_user_path);
static inline void file_free(struct file *f)
{
security_file_free(f);
if (likely(!(f->f_mode & FMODE_NOACCOUNT)))
percpu_counter_dec(&nr_files);
file: convert to SLAB_TYPESAFE_BY_RCU In recent discussions around some performance improvements in the file handling area we discussed switching the file cache to rely on SLAB_TYPESAFE_BY_RCU which allows us to get rid of call_rcu() based freeing for files completely. This is a pretty sensitive change overall but it might actually be worth doing. The main downside is the subtlety. The other one is that we should really wait for Jann's patch to land that enables KASAN to handle SLAB_TYPESAFE_BY_RCU UAFs. Currently it doesn't but a patch for this exists. With SLAB_TYPESAFE_BY_RCU objects may be freed and reused multiple times which requires a few changes. So it isn't sufficient anymore to just acquire a reference to the file in question under rcu using atomic_long_inc_not_zero() since the file might have already been recycled and someone else might have bumped the reference. In other words, callers might see reference count bumps from newer users. For this reason it is necessary to verify that the pointer is the same before and after the reference count increment. This pattern can be seen in get_file_rcu() and __files_get_rcu(). In addition, it isn't possible to access or check fields in struct file without first aqcuiring a reference on it. Not doing that was always very dodgy and it was only usable for non-pointer data in struct file. With SLAB_TYPESAFE_BY_RCU it is necessary that callers first acquire a reference under rcu or they must hold the files_lock of the fdtable. Failing to do either one of this is a bug. Thanks to Jann for pointing out that we need to ensure memory ordering between reallocations and pointer check by ensuring that all subsequent loads have a dependency on the second load in get_file_rcu() and providing a fixup that was folded into this patch. Cc: Jann Horn <jannh@google.com> Suggested-by: Linus Torvalds <torvalds@linux-foundation.org> Signed-off-by: Christian Brauner <brauner@kernel.org>
2023-09-29 06:45:59 +00:00
put_cred(f->f_cred);
if (unlikely(f->f_mode & FMODE_BACKING)) {
path_put(backing_file_user_path(f));
kmem_cache_free(bfilp_cachep, backing_file(f));
file: convert to SLAB_TYPESAFE_BY_RCU In recent discussions around some performance improvements in the file handling area we discussed switching the file cache to rely on SLAB_TYPESAFE_BY_RCU which allows us to get rid of call_rcu() based freeing for files completely. This is a pretty sensitive change overall but it might actually be worth doing. The main downside is the subtlety. The other one is that we should really wait for Jann's patch to land that enables KASAN to handle SLAB_TYPESAFE_BY_RCU UAFs. Currently it doesn't but a patch for this exists. With SLAB_TYPESAFE_BY_RCU objects may be freed and reused multiple times which requires a few changes. So it isn't sufficient anymore to just acquire a reference to the file in question under rcu using atomic_long_inc_not_zero() since the file might have already been recycled and someone else might have bumped the reference. In other words, callers might see reference count bumps from newer users. For this reason it is necessary to verify that the pointer is the same before and after the reference count increment. This pattern can be seen in get_file_rcu() and __files_get_rcu(). In addition, it isn't possible to access or check fields in struct file without first aqcuiring a reference on it. Not doing that was always very dodgy and it was only usable for non-pointer data in struct file. With SLAB_TYPESAFE_BY_RCU it is necessary that callers first acquire a reference under rcu or they must hold the files_lock of the fdtable. Failing to do either one of this is a bug. Thanks to Jann for pointing out that we need to ensure memory ordering between reallocations and pointer check by ensuring that all subsequent loads have a dependency on the second load in get_file_rcu() and providing a fixup that was folded into this patch. Cc: Jann Horn <jannh@google.com> Suggested-by: Linus Torvalds <torvalds@linux-foundation.org> Signed-off-by: Christian Brauner <brauner@kernel.org>
2023-09-29 06:45:59 +00:00
} else {
kmem_cache_free(filp_cachep, f);
}
}
/*
* Return the total number of open files in the system
*/
fs: allow for more than 2^31 files Robin Holt tried to boot a 16TB system and found af_unix was overflowing a 32bit value : <quote> We were seeing a failure which prevented boot. The kernel was incapable of creating either a named pipe or unix domain socket. This comes down to a common kernel function called unix_create1() which does: atomic_inc(&unix_nr_socks); if (atomic_read(&unix_nr_socks) > 2 * get_max_files()) goto out; The function get_max_files() is a simple return of files_stat.max_files. files_stat.max_files is a signed integer and is computed in fs/file_table.c's files_init(). n = (mempages * (PAGE_SIZE / 1024)) / 10; files_stat.max_files = n; In our case, mempages (total_ram_pages) is approx 3,758,096,384 (0xe0000000). That leaves max_files at approximately 1,503,238,553. This causes 2 * get_max_files() to integer overflow. </quote> Fix is to let /proc/sys/fs/file-nr & /proc/sys/fs/file-max use long integers, and change af_unix to use an atomic_long_t instead of atomic_t. get_max_files() is changed to return an unsigned long. get_nr_files() is changed to return a long. unix_nr_socks is changed from atomic_t to atomic_long_t, while not strictly needed to address Robin problem. Before patch (on a 64bit kernel) : # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max -18446744071562067968 After patch: # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max 2147483648 # cat /proc/sys/fs/file-nr 704 0 2147483648 Reported-by: Robin Holt <holt@sgi.com> Signed-off-by: Eric Dumazet <eric.dumazet@gmail.com> Acked-by: David Miller <davem@davemloft.net> Reviewed-by: Robin Holt <holt@sgi.com> Tested-by: Robin Holt <holt@sgi.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Christoph Hellwig <hch@lst.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-10-26 21:22:44 +00:00
static long get_nr_files(void)
{
return percpu_counter_read_positive(&nr_files);
}
/*
* Return the maximum number of open files in the system
*/
fs: allow for more than 2^31 files Robin Holt tried to boot a 16TB system and found af_unix was overflowing a 32bit value : <quote> We were seeing a failure which prevented boot. The kernel was incapable of creating either a named pipe or unix domain socket. This comes down to a common kernel function called unix_create1() which does: atomic_inc(&unix_nr_socks); if (atomic_read(&unix_nr_socks) > 2 * get_max_files()) goto out; The function get_max_files() is a simple return of files_stat.max_files. files_stat.max_files is a signed integer and is computed in fs/file_table.c's files_init(). n = (mempages * (PAGE_SIZE / 1024)) / 10; files_stat.max_files = n; In our case, mempages (total_ram_pages) is approx 3,758,096,384 (0xe0000000). That leaves max_files at approximately 1,503,238,553. This causes 2 * get_max_files() to integer overflow. </quote> Fix is to let /proc/sys/fs/file-nr & /proc/sys/fs/file-max use long integers, and change af_unix to use an atomic_long_t instead of atomic_t. get_max_files() is changed to return an unsigned long. get_nr_files() is changed to return a long. unix_nr_socks is changed from atomic_t to atomic_long_t, while not strictly needed to address Robin problem. Before patch (on a 64bit kernel) : # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max -18446744071562067968 After patch: # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max 2147483648 # cat /proc/sys/fs/file-nr 704 0 2147483648 Reported-by: Robin Holt <holt@sgi.com> Signed-off-by: Eric Dumazet <eric.dumazet@gmail.com> Acked-by: David Miller <davem@davemloft.net> Reviewed-by: Robin Holt <holt@sgi.com> Tested-by: Robin Holt <holt@sgi.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Christoph Hellwig <hch@lst.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-10-26 21:22:44 +00:00
unsigned long get_max_files(void)
{
return files_stat.max_files;
}
EXPORT_SYMBOL_GPL(get_max_files);
#if defined(CONFIG_SYSCTL) && defined(CONFIG_PROC_FS)
/*
* Handle nr_files sysctl
*/
sysctl: treewide: constify the ctl_table argument of proc_handlers const qualify the struct ctl_table argument in the proc_handler function signatures. This is a prerequisite to moving the static ctl_table structs into .rodata data which will ensure that proc_handler function pointers cannot be modified. This patch has been generated by the following coccinelle script: ``` virtual patch @r1@ identifier ctl, write, buffer, lenp, ppos; identifier func !~ "appldata_(timer|interval)_handler|sched_(rt|rr)_handler|rds_tcp_skbuf_handler|proc_sctp_do_(hmac_alg|rto_min|rto_max|udp_port|alpha_beta|auth|probe_interval)"; @@ int func( - struct ctl_table *ctl + const struct ctl_table *ctl ,int write, void *buffer, size_t *lenp, loff_t *ppos); @r2@ identifier func, ctl, write, buffer, lenp, ppos; @@ int func( - struct ctl_table *ctl + const struct ctl_table *ctl ,int write, void *buffer, size_t *lenp, loff_t *ppos) { ... } @r3@ identifier func; @@ int func( - struct ctl_table * + const struct ctl_table * ,int , void *, size_t *, loff_t *); @r4@ identifier func, ctl; @@ int func( - struct ctl_table *ctl + const struct ctl_table *ctl ,int , void *, size_t *, loff_t *); @r5@ identifier func, write, buffer, lenp, ppos; @@ int func( - struct ctl_table * + const struct ctl_table * ,int write, void *buffer, size_t *lenp, loff_t *ppos); ``` * Code formatting was adjusted in xfs_sysctl.c to comply with code conventions. The xfs_stats_clear_proc_handler, xfs_panic_mask_proc_handler and xfs_deprecated_dointvec_minmax where adjusted. * The ctl_table argument in proc_watchdog_common was const qualified. This is called from a proc_handler itself and is calling back into another proc_handler, making it necessary to change it as part of the proc_handler migration. Co-developed-by: Thomas Weißschuh <linux@weissschuh.net> Signed-off-by: Thomas Weißschuh <linux@weissschuh.net> Co-developed-by: Joel Granados <j.granados@samsung.com> Signed-off-by: Joel Granados <j.granados@samsung.com>
2024-07-24 18:59:29 +00:00
static int proc_nr_files(const struct ctl_table *table, int write, void *buffer,
size_t *lenp, loff_t *ppos)
{
files_stat.nr_files = get_nr_files();
fs: allow for more than 2^31 files Robin Holt tried to boot a 16TB system and found af_unix was overflowing a 32bit value : <quote> We were seeing a failure which prevented boot. The kernel was incapable of creating either a named pipe or unix domain socket. This comes down to a common kernel function called unix_create1() which does: atomic_inc(&unix_nr_socks); if (atomic_read(&unix_nr_socks) > 2 * get_max_files()) goto out; The function get_max_files() is a simple return of files_stat.max_files. files_stat.max_files is a signed integer and is computed in fs/file_table.c's files_init(). n = (mempages * (PAGE_SIZE / 1024)) / 10; files_stat.max_files = n; In our case, mempages (total_ram_pages) is approx 3,758,096,384 (0xe0000000). That leaves max_files at approximately 1,503,238,553. This causes 2 * get_max_files() to integer overflow. </quote> Fix is to let /proc/sys/fs/file-nr & /proc/sys/fs/file-max use long integers, and change af_unix to use an atomic_long_t instead of atomic_t. get_max_files() is changed to return an unsigned long. get_nr_files() is changed to return a long. unix_nr_socks is changed from atomic_t to atomic_long_t, while not strictly needed to address Robin problem. Before patch (on a 64bit kernel) : # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max -18446744071562067968 After patch: # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max 2147483648 # cat /proc/sys/fs/file-nr 704 0 2147483648 Reported-by: Robin Holt <holt@sgi.com> Signed-off-by: Eric Dumazet <eric.dumazet@gmail.com> Acked-by: David Miller <davem@davemloft.net> Reviewed-by: Robin Holt <holt@sgi.com> Tested-by: Robin Holt <holt@sgi.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Christoph Hellwig <hch@lst.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-10-26 21:22:44 +00:00
return proc_doulongvec_minmax(table, write, buffer, lenp, ppos);
}
static struct ctl_table fs_stat_sysctls[] = {
{
.procname = "file-nr",
.data = &files_stat,
.maxlen = sizeof(files_stat),
.mode = 0444,
.proc_handler = proc_nr_files,
},
{
.procname = "file-max",
.data = &files_stat.max_files,
.maxlen = sizeof(files_stat.max_files),
.mode = 0644,
.proc_handler = proc_doulongvec_minmax,
.extra1 = SYSCTL_LONG_ZERO,
.extra2 = SYSCTL_LONG_MAX,
},
{
.procname = "nr_open",
.data = &sysctl_nr_open,
.maxlen = sizeof(unsigned int),
.mode = 0644,
.proc_handler = proc_dointvec_minmax,
.extra1 = &sysctl_nr_open_min,
.extra2 = &sysctl_nr_open_max,
},
};
static int __init init_fs_stat_sysctls(void)
{
register_sysctl_init("fs", fs_stat_sysctls);
if (IS_ENABLED(CONFIG_BINFMT_MISC)) {
struct ctl_table_header *hdr;
hdr = register_sysctl_mount_point("fs/binfmt_misc");
kmemleak_not_leak(hdr);
}
return 0;
}
fs_initcall(init_fs_stat_sysctls);
#endif
static int init_file(struct file *f, int flags, const struct cred *cred)
{
int error;
f->f_cred = get_cred(cred);
error = security_file_alloc(f);
if (unlikely(error)) {
put_cred(f->f_cred);
return error;
}
spin_lock_init(&f->f_lock);
/*
* Note that f_pos_lock is only used for files raising
* FMODE_ATOMIC_POS and directories. Other files such as pipes
* don't need it and since f_pos_lock is in a union may reuse
* the space for other purposes. They are expected to initialize
* the respective member when opening the file.
*/
mutex_init(&f->f_pos_lock);
fs: port files to file_ref Port files to rely on file_ref reference to improve scaling and gain overflow protection. - We continue to WARN during get_file() in case a file that is already marked dead is revived as get_file() is only valid if the caller already holds a reference to the file. This hasn't changed just the check changes. - The semantics for epoll and ttm's dmabuf usage have changed. Both epoll and ttm synchronize with __fput() to prevent the underlying file from beeing freed. (1) epoll Explaining epoll is straightforward using a simple diagram. Essentially, the mutex of the epoll instance needs to be taken in both __fput() and around epi_fget() preventing the file from being freed while it is polled or preventing the file from being resurrected. CPU1 CPU2 fput(file) -> __fput(file) -> eventpoll_release(file) -> eventpoll_release_file(file) mutex_lock(&ep->mtx) epi_item_poll() -> epi_fget() -> file_ref_get(file) mutex_unlock(&ep->mtx) mutex_lock(&ep->mtx); __ep_remove() mutex_unlock(&ep->mtx); -> kmem_cache_free(file) (2) ttm dmabuf This explanation is a bit more involved. A regular dmabuf file stashed the dmabuf in file->private_data and the file in dmabuf->file: file->private_data = dmabuf; dmabuf->file = file; The generic release method of a dmabuf file handles file specific things: f_op->release::dma_buf_file_release() while the generic dentry release method of a dmabuf handles dmabuf freeing including driver specific things: dentry->d_release::dma_buf_release() During ttm dmabuf initialization in ttm_object_device_init() the ttm driver copies the provided struct dma_buf_ops into a private location: struct ttm_object_device { spinlock_t object_lock; struct dma_buf_ops ops; void (*dmabuf_release)(struct dma_buf *dma_buf); struct idr idr; }; ttm_object_device_init(const struct dma_buf_ops *ops) { // copy original dma_buf_ops in private location tdev->ops = *ops; // stash the release method of the original struct dma_buf_ops tdev->dmabuf_release = tdev->ops.release; // override the release method in the copy of the struct dma_buf_ops // with ttm's own dmabuf release method tdev->ops.release = ttm_prime_dmabuf_release; } When a new dmabuf is created the struct dma_buf_ops with the overriden release method set to ttm_prime_dmabuf_release is passed in exp_info.ops: DEFINE_DMA_BUF_EXPORT_INFO(exp_info); exp_info.ops = &tdev->ops; exp_info.size = prime->size; exp_info.flags = flags; exp_info.priv = prime; The call to dma_buf_export() then sets mutex_lock_interruptible(&prime->mutex); dma_buf = dma_buf_export(&exp_info) { dmabuf->ops = exp_info->ops; } mutex_unlock(&prime->mutex); which creates a new dmabuf file and then install a file descriptor to it in the callers file descriptor table: ret = dma_buf_fd(dma_buf, flags); When that dmabuf file is closed we now get: fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); Where we can see that prime->dma_buf is set to NULL. So when we have the following diagram: CPU1 CPU2 fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() ttm_prime_handle_to_fd() mutex_lock_interruptible(&prime->mutex) dma_buf = prime->dma_buf dma_buf && get_dma_buf_unless_doomed(dma_buf) -> file_ref_get(dma_buf->file) mutex_unlock(&prime->mutex); mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); -> kmem_cache_free(file) The logic of the mechanism is the same as for epoll: sync with __fput() preventing the file from being freed. Here the synchronization happens through the ttm instance's prime->mutex. Basically, the lifetime of the dma_buf and the file are tighly coupled. Both (1) and (2) used to call atomic_inc_not_zero() to check whether the file has already been marked dead and then refuse to revive it. This is only safe because both (1) and (2) sync with __fput() and thus prevent kmem_cache_free() on the file being called and thus prevent the file from being immediately recycled due to SLAB_TYPESAFE_BY_RCU. Both (1) and (2) have been ported from atomic_inc_not_zero() to file_ref_get(). That means a file that is already in the process of being marked as FILE_REF_DEAD: file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) can be revived again: CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) This is fine and inherent to the file_ref_get()/file_ref_put() semantics. For both (1) and (2) this is safe because __fput() is prevented from making progress if file_ref_get() fails due to the aforementioned synchronization mechanisms. Two cases need to be considered that affect both (1) epoll and (2) ttm dmabuf: (i) fput()'s file_ref_put() and marks the file as FILE_REF_NOREF but before that fput() can mark the file as FILE_REF_DEAD someone manages to sneak in a file_ref_get() and brings the refcount back from FILE_REF_NOREF to FILE_REF_ONEREF. In that case the original fput() doesn't call __fput(). For epoll the poll will finish and for ttm dmabuf the file can be used again. For ttm dambuf this is actually an advantage because it avoids immediately allocating a new dmabuf object. CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) (ii) fput()'s file_ref_put() marks the file FILE_REF_NOREF and also suceeds in actually marking it FILE_REF_DEAD and then calls into __fput() to free the file. When either (1) or (2) call file_ref_get() they fail as atomic_long_add_negative() will return true. At the same time, both (1) and (2) all file_ref_get() under mutexes that __fput() must also acquire preventing kmem_cache_free() from freeing the file. So while this might be treated as a change in semantics for (1) and (2) it really isn't. It if should end up causing issues this can be fixed by adding a helper that does something like: long cnt = atomic_long_read(&ref->refcnt); do { if (cnt < 0) return false; } while (!atomic_long_try_cmpxchg(&ref->refcnt, &cnt, cnt + 1)); return true; which would block FILE_REF_NOREF to FILE_REF_ONEREF transitions. - Jann correctly pointed out that kmem_cache_zalloc() cannot be used anymore once files have been ported to file_ref_t. The kmem_cache_zalloc() call will memset() the whole struct file to zero when it is reallocated. This will also set file->f_ref to zero which mens that a concurrent file_ref_get() can return true: CPU1 CPU2 __get_file_rcu() rcu_dereference_raw() close() [frees file] alloc_empty_file() kmem_cache_zalloc() [reallocates same file] memset(..., 0, ...) file_ref_get() [increments 0->1, returns true] init_file() file_ref_init(..., 1) [sets to 0] rcu_dereference_raw() fput() file_ref_put() [decrements 0->FILE_REF_NOREF, frees file] [UAF] causing a concurrent __get_file_rcu() call to acquire a reference to the file that is about to be reallocated and immediately freeing it on realizing that it has been recycled. This causes a UAF for the task that reallocated/recycled the file. This is prevented by switching from kmem_cache_zalloc() to kmem_cache_alloc() and initializing the fields manually. With file->f_ref initialized last. Note that a memset() also isn't guaranteed to atomically update an unsigned long so it's theoretically possible to see torn and therefore bogus counter values. Link: https://lore.kernel.org/r/20241007-brauner-file-rcuref-v2-3-387e24dc9163@kernel.org Signed-off-by: Christian Brauner <brauner@kernel.org>
2024-10-07 14:23:59 +00:00
memset(&f->f_path, 0, sizeof(f->f_path));
memset(&f->f_ra, 0, sizeof(f->f_ra));
f->f_flags = flags;
f->f_mode = OPEN_FMODE(flags);
f->f_op = NULL;
f->f_mapping = NULL;
f->private_data = NULL;
f->f_inode = NULL;
f->f_owner = NULL;
#ifdef CONFIG_EPOLL
f->f_ep = NULL;
#endif
f->f_iocb_flags = 0;
f->f_pos = 0;
f->f_wb_err = 0;
f->f_sb_err = 0;
file: convert to SLAB_TYPESAFE_BY_RCU In recent discussions around some performance improvements in the file handling area we discussed switching the file cache to rely on SLAB_TYPESAFE_BY_RCU which allows us to get rid of call_rcu() based freeing for files completely. This is a pretty sensitive change overall but it might actually be worth doing. The main downside is the subtlety. The other one is that we should really wait for Jann's patch to land that enables KASAN to handle SLAB_TYPESAFE_BY_RCU UAFs. Currently it doesn't but a patch for this exists. With SLAB_TYPESAFE_BY_RCU objects may be freed and reused multiple times which requires a few changes. So it isn't sufficient anymore to just acquire a reference to the file in question under rcu using atomic_long_inc_not_zero() since the file might have already been recycled and someone else might have bumped the reference. In other words, callers might see reference count bumps from newer users. For this reason it is necessary to verify that the pointer is the same before and after the reference count increment. This pattern can be seen in get_file_rcu() and __files_get_rcu(). In addition, it isn't possible to access or check fields in struct file without first aqcuiring a reference on it. Not doing that was always very dodgy and it was only usable for non-pointer data in struct file. With SLAB_TYPESAFE_BY_RCU it is necessary that callers first acquire a reference under rcu or they must hold the files_lock of the fdtable. Failing to do either one of this is a bug. Thanks to Jann for pointing out that we need to ensure memory ordering between reallocations and pointer check by ensuring that all subsequent loads have a dependency on the second load in get_file_rcu() and providing a fixup that was folded into this patch. Cc: Jann Horn <jannh@google.com> Suggested-by: Linus Torvalds <torvalds@linux-foundation.org> Signed-off-by: Christian Brauner <brauner@kernel.org>
2023-09-29 06:45:59 +00:00
/*
* We're SLAB_TYPESAFE_BY_RCU so initialize f_count last. While
* fget-rcu pattern users need to be able to handle spurious
* refcount bumps we should reinitialize the reused file first.
*/
fs: port files to file_ref Port files to rely on file_ref reference to improve scaling and gain overflow protection. - We continue to WARN during get_file() in case a file that is already marked dead is revived as get_file() is only valid if the caller already holds a reference to the file. This hasn't changed just the check changes. - The semantics for epoll and ttm's dmabuf usage have changed. Both epoll and ttm synchronize with __fput() to prevent the underlying file from beeing freed. (1) epoll Explaining epoll is straightforward using a simple diagram. Essentially, the mutex of the epoll instance needs to be taken in both __fput() and around epi_fget() preventing the file from being freed while it is polled or preventing the file from being resurrected. CPU1 CPU2 fput(file) -> __fput(file) -> eventpoll_release(file) -> eventpoll_release_file(file) mutex_lock(&ep->mtx) epi_item_poll() -> epi_fget() -> file_ref_get(file) mutex_unlock(&ep->mtx) mutex_lock(&ep->mtx); __ep_remove() mutex_unlock(&ep->mtx); -> kmem_cache_free(file) (2) ttm dmabuf This explanation is a bit more involved. A regular dmabuf file stashed the dmabuf in file->private_data and the file in dmabuf->file: file->private_data = dmabuf; dmabuf->file = file; The generic release method of a dmabuf file handles file specific things: f_op->release::dma_buf_file_release() while the generic dentry release method of a dmabuf handles dmabuf freeing including driver specific things: dentry->d_release::dma_buf_release() During ttm dmabuf initialization in ttm_object_device_init() the ttm driver copies the provided struct dma_buf_ops into a private location: struct ttm_object_device { spinlock_t object_lock; struct dma_buf_ops ops; void (*dmabuf_release)(struct dma_buf *dma_buf); struct idr idr; }; ttm_object_device_init(const struct dma_buf_ops *ops) { // copy original dma_buf_ops in private location tdev->ops = *ops; // stash the release method of the original struct dma_buf_ops tdev->dmabuf_release = tdev->ops.release; // override the release method in the copy of the struct dma_buf_ops // with ttm's own dmabuf release method tdev->ops.release = ttm_prime_dmabuf_release; } When a new dmabuf is created the struct dma_buf_ops with the overriden release method set to ttm_prime_dmabuf_release is passed in exp_info.ops: DEFINE_DMA_BUF_EXPORT_INFO(exp_info); exp_info.ops = &tdev->ops; exp_info.size = prime->size; exp_info.flags = flags; exp_info.priv = prime; The call to dma_buf_export() then sets mutex_lock_interruptible(&prime->mutex); dma_buf = dma_buf_export(&exp_info) { dmabuf->ops = exp_info->ops; } mutex_unlock(&prime->mutex); which creates a new dmabuf file and then install a file descriptor to it in the callers file descriptor table: ret = dma_buf_fd(dma_buf, flags); When that dmabuf file is closed we now get: fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); Where we can see that prime->dma_buf is set to NULL. So when we have the following diagram: CPU1 CPU2 fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() ttm_prime_handle_to_fd() mutex_lock_interruptible(&prime->mutex) dma_buf = prime->dma_buf dma_buf && get_dma_buf_unless_doomed(dma_buf) -> file_ref_get(dma_buf->file) mutex_unlock(&prime->mutex); mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); -> kmem_cache_free(file) The logic of the mechanism is the same as for epoll: sync with __fput() preventing the file from being freed. Here the synchronization happens through the ttm instance's prime->mutex. Basically, the lifetime of the dma_buf and the file are tighly coupled. Both (1) and (2) used to call atomic_inc_not_zero() to check whether the file has already been marked dead and then refuse to revive it. This is only safe because both (1) and (2) sync with __fput() and thus prevent kmem_cache_free() on the file being called and thus prevent the file from being immediately recycled due to SLAB_TYPESAFE_BY_RCU. Both (1) and (2) have been ported from atomic_inc_not_zero() to file_ref_get(). That means a file that is already in the process of being marked as FILE_REF_DEAD: file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) can be revived again: CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) This is fine and inherent to the file_ref_get()/file_ref_put() semantics. For both (1) and (2) this is safe because __fput() is prevented from making progress if file_ref_get() fails due to the aforementioned synchronization mechanisms. Two cases need to be considered that affect both (1) epoll and (2) ttm dmabuf: (i) fput()'s file_ref_put() and marks the file as FILE_REF_NOREF but before that fput() can mark the file as FILE_REF_DEAD someone manages to sneak in a file_ref_get() and brings the refcount back from FILE_REF_NOREF to FILE_REF_ONEREF. In that case the original fput() doesn't call __fput(). For epoll the poll will finish and for ttm dmabuf the file can be used again. For ttm dambuf this is actually an advantage because it avoids immediately allocating a new dmabuf object. CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) (ii) fput()'s file_ref_put() marks the file FILE_REF_NOREF and also suceeds in actually marking it FILE_REF_DEAD and then calls into __fput() to free the file. When either (1) or (2) call file_ref_get() they fail as atomic_long_add_negative() will return true. At the same time, both (1) and (2) all file_ref_get() under mutexes that __fput() must also acquire preventing kmem_cache_free() from freeing the file. So while this might be treated as a change in semantics for (1) and (2) it really isn't. It if should end up causing issues this can be fixed by adding a helper that does something like: long cnt = atomic_long_read(&ref->refcnt); do { if (cnt < 0) return false; } while (!atomic_long_try_cmpxchg(&ref->refcnt, &cnt, cnt + 1)); return true; which would block FILE_REF_NOREF to FILE_REF_ONEREF transitions. - Jann correctly pointed out that kmem_cache_zalloc() cannot be used anymore once files have been ported to file_ref_t. The kmem_cache_zalloc() call will memset() the whole struct file to zero when it is reallocated. This will also set file->f_ref to zero which mens that a concurrent file_ref_get() can return true: CPU1 CPU2 __get_file_rcu() rcu_dereference_raw() close() [frees file] alloc_empty_file() kmem_cache_zalloc() [reallocates same file] memset(..., 0, ...) file_ref_get() [increments 0->1, returns true] init_file() file_ref_init(..., 1) [sets to 0] rcu_dereference_raw() fput() file_ref_put() [decrements 0->FILE_REF_NOREF, frees file] [UAF] causing a concurrent __get_file_rcu() call to acquire a reference to the file that is about to be reallocated and immediately freeing it on realizing that it has been recycled. This causes a UAF for the task that reallocated/recycled the file. This is prevented by switching from kmem_cache_zalloc() to kmem_cache_alloc() and initializing the fields manually. With file->f_ref initialized last. Note that a memset() also isn't guaranteed to atomically update an unsigned long so it's theoretically possible to see torn and therefore bogus counter values. Link: https://lore.kernel.org/r/20241007-brauner-file-rcuref-v2-3-387e24dc9163@kernel.org Signed-off-by: Christian Brauner <brauner@kernel.org>
2024-10-07 14:23:59 +00:00
file_ref_init(&f->f_ref, 1);
return 0;
}
/* Find an unused file structure and return a pointer to it.
* Returns an error pointer if some error happend e.g. we over file
* structures limit, run out of memory or operation is not permitted.
*
* Be very careful using this. You are responsible for
* getting write access to any mount that you might assign
* to this filp, if it is opened for write. If this is not
* done, you will imbalance int the mount's writer count
* and a warning at __fput() time.
*/
struct file *alloc_empty_file(int flags, const struct cred *cred)
{
fs: allow for more than 2^31 files Robin Holt tried to boot a 16TB system and found af_unix was overflowing a 32bit value : <quote> We were seeing a failure which prevented boot. The kernel was incapable of creating either a named pipe or unix domain socket. This comes down to a common kernel function called unix_create1() which does: atomic_inc(&unix_nr_socks); if (atomic_read(&unix_nr_socks) > 2 * get_max_files()) goto out; The function get_max_files() is a simple return of files_stat.max_files. files_stat.max_files is a signed integer and is computed in fs/file_table.c's files_init(). n = (mempages * (PAGE_SIZE / 1024)) / 10; files_stat.max_files = n; In our case, mempages (total_ram_pages) is approx 3,758,096,384 (0xe0000000). That leaves max_files at approximately 1,503,238,553. This causes 2 * get_max_files() to integer overflow. </quote> Fix is to let /proc/sys/fs/file-nr & /proc/sys/fs/file-max use long integers, and change af_unix to use an atomic_long_t instead of atomic_t. get_max_files() is changed to return an unsigned long. get_nr_files() is changed to return a long. unix_nr_socks is changed from atomic_t to atomic_long_t, while not strictly needed to address Robin problem. Before patch (on a 64bit kernel) : # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max -18446744071562067968 After patch: # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max 2147483648 # cat /proc/sys/fs/file-nr 704 0 2147483648 Reported-by: Robin Holt <holt@sgi.com> Signed-off-by: Eric Dumazet <eric.dumazet@gmail.com> Acked-by: David Miller <davem@davemloft.net> Reviewed-by: Robin Holt <holt@sgi.com> Tested-by: Robin Holt <holt@sgi.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Christoph Hellwig <hch@lst.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-10-26 21:22:44 +00:00
static long old_max;
struct file *f;
int error;
/*
* Privileged users can go above max_files
*/
if (get_nr_files() >= files_stat.max_files && !capable(CAP_SYS_ADMIN)) {
/*
* percpu_counters are inaccurate. Do an expensive check before
* we go and fail.
*/
if (percpu_counter_sum_positive(&nr_files) >= files_stat.max_files)
goto over;
}
fs: port files to file_ref Port files to rely on file_ref reference to improve scaling and gain overflow protection. - We continue to WARN during get_file() in case a file that is already marked dead is revived as get_file() is only valid if the caller already holds a reference to the file. This hasn't changed just the check changes. - The semantics for epoll and ttm's dmabuf usage have changed. Both epoll and ttm synchronize with __fput() to prevent the underlying file from beeing freed. (1) epoll Explaining epoll is straightforward using a simple diagram. Essentially, the mutex of the epoll instance needs to be taken in both __fput() and around epi_fget() preventing the file from being freed while it is polled or preventing the file from being resurrected. CPU1 CPU2 fput(file) -> __fput(file) -> eventpoll_release(file) -> eventpoll_release_file(file) mutex_lock(&ep->mtx) epi_item_poll() -> epi_fget() -> file_ref_get(file) mutex_unlock(&ep->mtx) mutex_lock(&ep->mtx); __ep_remove() mutex_unlock(&ep->mtx); -> kmem_cache_free(file) (2) ttm dmabuf This explanation is a bit more involved. A regular dmabuf file stashed the dmabuf in file->private_data and the file in dmabuf->file: file->private_data = dmabuf; dmabuf->file = file; The generic release method of a dmabuf file handles file specific things: f_op->release::dma_buf_file_release() while the generic dentry release method of a dmabuf handles dmabuf freeing including driver specific things: dentry->d_release::dma_buf_release() During ttm dmabuf initialization in ttm_object_device_init() the ttm driver copies the provided struct dma_buf_ops into a private location: struct ttm_object_device { spinlock_t object_lock; struct dma_buf_ops ops; void (*dmabuf_release)(struct dma_buf *dma_buf); struct idr idr; }; ttm_object_device_init(const struct dma_buf_ops *ops) { // copy original dma_buf_ops in private location tdev->ops = *ops; // stash the release method of the original struct dma_buf_ops tdev->dmabuf_release = tdev->ops.release; // override the release method in the copy of the struct dma_buf_ops // with ttm's own dmabuf release method tdev->ops.release = ttm_prime_dmabuf_release; } When a new dmabuf is created the struct dma_buf_ops with the overriden release method set to ttm_prime_dmabuf_release is passed in exp_info.ops: DEFINE_DMA_BUF_EXPORT_INFO(exp_info); exp_info.ops = &tdev->ops; exp_info.size = prime->size; exp_info.flags = flags; exp_info.priv = prime; The call to dma_buf_export() then sets mutex_lock_interruptible(&prime->mutex); dma_buf = dma_buf_export(&exp_info) { dmabuf->ops = exp_info->ops; } mutex_unlock(&prime->mutex); which creates a new dmabuf file and then install a file descriptor to it in the callers file descriptor table: ret = dma_buf_fd(dma_buf, flags); When that dmabuf file is closed we now get: fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); Where we can see that prime->dma_buf is set to NULL. So when we have the following diagram: CPU1 CPU2 fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() ttm_prime_handle_to_fd() mutex_lock_interruptible(&prime->mutex) dma_buf = prime->dma_buf dma_buf && get_dma_buf_unless_doomed(dma_buf) -> file_ref_get(dma_buf->file) mutex_unlock(&prime->mutex); mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); -> kmem_cache_free(file) The logic of the mechanism is the same as for epoll: sync with __fput() preventing the file from being freed. Here the synchronization happens through the ttm instance's prime->mutex. Basically, the lifetime of the dma_buf and the file are tighly coupled. Both (1) and (2) used to call atomic_inc_not_zero() to check whether the file has already been marked dead and then refuse to revive it. This is only safe because both (1) and (2) sync with __fput() and thus prevent kmem_cache_free() on the file being called and thus prevent the file from being immediately recycled due to SLAB_TYPESAFE_BY_RCU. Both (1) and (2) have been ported from atomic_inc_not_zero() to file_ref_get(). That means a file that is already in the process of being marked as FILE_REF_DEAD: file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) can be revived again: CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) This is fine and inherent to the file_ref_get()/file_ref_put() semantics. For both (1) and (2) this is safe because __fput() is prevented from making progress if file_ref_get() fails due to the aforementioned synchronization mechanisms. Two cases need to be considered that affect both (1) epoll and (2) ttm dmabuf: (i) fput()'s file_ref_put() and marks the file as FILE_REF_NOREF but before that fput() can mark the file as FILE_REF_DEAD someone manages to sneak in a file_ref_get() and brings the refcount back from FILE_REF_NOREF to FILE_REF_ONEREF. In that case the original fput() doesn't call __fput(). For epoll the poll will finish and for ttm dmabuf the file can be used again. For ttm dambuf this is actually an advantage because it avoids immediately allocating a new dmabuf object. CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) (ii) fput()'s file_ref_put() marks the file FILE_REF_NOREF and also suceeds in actually marking it FILE_REF_DEAD and then calls into __fput() to free the file. When either (1) or (2) call file_ref_get() they fail as atomic_long_add_negative() will return true. At the same time, both (1) and (2) all file_ref_get() under mutexes that __fput() must also acquire preventing kmem_cache_free() from freeing the file. So while this might be treated as a change in semantics for (1) and (2) it really isn't. It if should end up causing issues this can be fixed by adding a helper that does something like: long cnt = atomic_long_read(&ref->refcnt); do { if (cnt < 0) return false; } while (!atomic_long_try_cmpxchg(&ref->refcnt, &cnt, cnt + 1)); return true; which would block FILE_REF_NOREF to FILE_REF_ONEREF transitions. - Jann correctly pointed out that kmem_cache_zalloc() cannot be used anymore once files have been ported to file_ref_t. The kmem_cache_zalloc() call will memset() the whole struct file to zero when it is reallocated. This will also set file->f_ref to zero which mens that a concurrent file_ref_get() can return true: CPU1 CPU2 __get_file_rcu() rcu_dereference_raw() close() [frees file] alloc_empty_file() kmem_cache_zalloc() [reallocates same file] memset(..., 0, ...) file_ref_get() [increments 0->1, returns true] init_file() file_ref_init(..., 1) [sets to 0] rcu_dereference_raw() fput() file_ref_put() [decrements 0->FILE_REF_NOREF, frees file] [UAF] causing a concurrent __get_file_rcu() call to acquire a reference to the file that is about to be reallocated and immediately freeing it on realizing that it has been recycled. This causes a UAF for the task that reallocated/recycled the file. This is prevented by switching from kmem_cache_zalloc() to kmem_cache_alloc() and initializing the fields manually. With file->f_ref initialized last. Note that a memset() also isn't guaranteed to atomically update an unsigned long so it's theoretically possible to see torn and therefore bogus counter values. Link: https://lore.kernel.org/r/20241007-brauner-file-rcuref-v2-3-387e24dc9163@kernel.org Signed-off-by: Christian Brauner <brauner@kernel.org>
2024-10-07 14:23:59 +00:00
f = kmem_cache_alloc(filp_cachep, GFP_KERNEL);
if (unlikely(!f))
return ERR_PTR(-ENOMEM);
error = init_file(f, flags, cred);
if (unlikely(error)) {
kmem_cache_free(filp_cachep, f);
return ERR_PTR(error);
}
percpu_counter_inc(&nr_files);
return f;
over:
/* Ran out of filps - report that */
if (get_nr_files() > old_max) {
fs: allow for more than 2^31 files Robin Holt tried to boot a 16TB system and found af_unix was overflowing a 32bit value : <quote> We were seeing a failure which prevented boot. The kernel was incapable of creating either a named pipe or unix domain socket. This comes down to a common kernel function called unix_create1() which does: atomic_inc(&unix_nr_socks); if (atomic_read(&unix_nr_socks) > 2 * get_max_files()) goto out; The function get_max_files() is a simple return of files_stat.max_files. files_stat.max_files is a signed integer and is computed in fs/file_table.c's files_init(). n = (mempages * (PAGE_SIZE / 1024)) / 10; files_stat.max_files = n; In our case, mempages (total_ram_pages) is approx 3,758,096,384 (0xe0000000). That leaves max_files at approximately 1,503,238,553. This causes 2 * get_max_files() to integer overflow. </quote> Fix is to let /proc/sys/fs/file-nr & /proc/sys/fs/file-max use long integers, and change af_unix to use an atomic_long_t instead of atomic_t. get_max_files() is changed to return an unsigned long. get_nr_files() is changed to return a long. unix_nr_socks is changed from atomic_t to atomic_long_t, while not strictly needed to address Robin problem. Before patch (on a 64bit kernel) : # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max -18446744071562067968 After patch: # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max 2147483648 # cat /proc/sys/fs/file-nr 704 0 2147483648 Reported-by: Robin Holt <holt@sgi.com> Signed-off-by: Eric Dumazet <eric.dumazet@gmail.com> Acked-by: David Miller <davem@davemloft.net> Reviewed-by: Robin Holt <holt@sgi.com> Tested-by: Robin Holt <holt@sgi.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Christoph Hellwig <hch@lst.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-10-26 21:22:44 +00:00
pr_info("VFS: file-max limit %lu reached\n", get_max_files());
old_max = get_nr_files();
}
return ERR_PTR(-ENFILE);
}
/*
* Variant of alloc_empty_file() that doesn't check and modify nr_files.
*
* This is only for kernel internal use, and the allocate file must not be
* installed into file tables or such.
*/
struct file *alloc_empty_file_noaccount(int flags, const struct cred *cred)
{
struct file *f;
int error;
fs: port files to file_ref Port files to rely on file_ref reference to improve scaling and gain overflow protection. - We continue to WARN during get_file() in case a file that is already marked dead is revived as get_file() is only valid if the caller already holds a reference to the file. This hasn't changed just the check changes. - The semantics for epoll and ttm's dmabuf usage have changed. Both epoll and ttm synchronize with __fput() to prevent the underlying file from beeing freed. (1) epoll Explaining epoll is straightforward using a simple diagram. Essentially, the mutex of the epoll instance needs to be taken in both __fput() and around epi_fget() preventing the file from being freed while it is polled or preventing the file from being resurrected. CPU1 CPU2 fput(file) -> __fput(file) -> eventpoll_release(file) -> eventpoll_release_file(file) mutex_lock(&ep->mtx) epi_item_poll() -> epi_fget() -> file_ref_get(file) mutex_unlock(&ep->mtx) mutex_lock(&ep->mtx); __ep_remove() mutex_unlock(&ep->mtx); -> kmem_cache_free(file) (2) ttm dmabuf This explanation is a bit more involved. A regular dmabuf file stashed the dmabuf in file->private_data and the file in dmabuf->file: file->private_data = dmabuf; dmabuf->file = file; The generic release method of a dmabuf file handles file specific things: f_op->release::dma_buf_file_release() while the generic dentry release method of a dmabuf handles dmabuf freeing including driver specific things: dentry->d_release::dma_buf_release() During ttm dmabuf initialization in ttm_object_device_init() the ttm driver copies the provided struct dma_buf_ops into a private location: struct ttm_object_device { spinlock_t object_lock; struct dma_buf_ops ops; void (*dmabuf_release)(struct dma_buf *dma_buf); struct idr idr; }; ttm_object_device_init(const struct dma_buf_ops *ops) { // copy original dma_buf_ops in private location tdev->ops = *ops; // stash the release method of the original struct dma_buf_ops tdev->dmabuf_release = tdev->ops.release; // override the release method in the copy of the struct dma_buf_ops // with ttm's own dmabuf release method tdev->ops.release = ttm_prime_dmabuf_release; } When a new dmabuf is created the struct dma_buf_ops with the overriden release method set to ttm_prime_dmabuf_release is passed in exp_info.ops: DEFINE_DMA_BUF_EXPORT_INFO(exp_info); exp_info.ops = &tdev->ops; exp_info.size = prime->size; exp_info.flags = flags; exp_info.priv = prime; The call to dma_buf_export() then sets mutex_lock_interruptible(&prime->mutex); dma_buf = dma_buf_export(&exp_info) { dmabuf->ops = exp_info->ops; } mutex_unlock(&prime->mutex); which creates a new dmabuf file and then install a file descriptor to it in the callers file descriptor table: ret = dma_buf_fd(dma_buf, flags); When that dmabuf file is closed we now get: fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); Where we can see that prime->dma_buf is set to NULL. So when we have the following diagram: CPU1 CPU2 fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() ttm_prime_handle_to_fd() mutex_lock_interruptible(&prime->mutex) dma_buf = prime->dma_buf dma_buf && get_dma_buf_unless_doomed(dma_buf) -> file_ref_get(dma_buf->file) mutex_unlock(&prime->mutex); mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); -> kmem_cache_free(file) The logic of the mechanism is the same as for epoll: sync with __fput() preventing the file from being freed. Here the synchronization happens through the ttm instance's prime->mutex. Basically, the lifetime of the dma_buf and the file are tighly coupled. Both (1) and (2) used to call atomic_inc_not_zero() to check whether the file has already been marked dead and then refuse to revive it. This is only safe because both (1) and (2) sync with __fput() and thus prevent kmem_cache_free() on the file being called and thus prevent the file from being immediately recycled due to SLAB_TYPESAFE_BY_RCU. Both (1) and (2) have been ported from atomic_inc_not_zero() to file_ref_get(). That means a file that is already in the process of being marked as FILE_REF_DEAD: file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) can be revived again: CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) This is fine and inherent to the file_ref_get()/file_ref_put() semantics. For both (1) and (2) this is safe because __fput() is prevented from making progress if file_ref_get() fails due to the aforementioned synchronization mechanisms. Two cases need to be considered that affect both (1) epoll and (2) ttm dmabuf: (i) fput()'s file_ref_put() and marks the file as FILE_REF_NOREF but before that fput() can mark the file as FILE_REF_DEAD someone manages to sneak in a file_ref_get() and brings the refcount back from FILE_REF_NOREF to FILE_REF_ONEREF. In that case the original fput() doesn't call __fput(). For epoll the poll will finish and for ttm dmabuf the file can be used again. For ttm dambuf this is actually an advantage because it avoids immediately allocating a new dmabuf object. CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) (ii) fput()'s file_ref_put() marks the file FILE_REF_NOREF and also suceeds in actually marking it FILE_REF_DEAD and then calls into __fput() to free the file. When either (1) or (2) call file_ref_get() they fail as atomic_long_add_negative() will return true. At the same time, both (1) and (2) all file_ref_get() under mutexes that __fput() must also acquire preventing kmem_cache_free() from freeing the file. So while this might be treated as a change in semantics for (1) and (2) it really isn't. It if should end up causing issues this can be fixed by adding a helper that does something like: long cnt = atomic_long_read(&ref->refcnt); do { if (cnt < 0) return false; } while (!atomic_long_try_cmpxchg(&ref->refcnt, &cnt, cnt + 1)); return true; which would block FILE_REF_NOREF to FILE_REF_ONEREF transitions. - Jann correctly pointed out that kmem_cache_zalloc() cannot be used anymore once files have been ported to file_ref_t. The kmem_cache_zalloc() call will memset() the whole struct file to zero when it is reallocated. This will also set file->f_ref to zero which mens that a concurrent file_ref_get() can return true: CPU1 CPU2 __get_file_rcu() rcu_dereference_raw() close() [frees file] alloc_empty_file() kmem_cache_zalloc() [reallocates same file] memset(..., 0, ...) file_ref_get() [increments 0->1, returns true] init_file() file_ref_init(..., 1) [sets to 0] rcu_dereference_raw() fput() file_ref_put() [decrements 0->FILE_REF_NOREF, frees file] [UAF] causing a concurrent __get_file_rcu() call to acquire a reference to the file that is about to be reallocated and immediately freeing it on realizing that it has been recycled. This causes a UAF for the task that reallocated/recycled the file. This is prevented by switching from kmem_cache_zalloc() to kmem_cache_alloc() and initializing the fields manually. With file->f_ref initialized last. Note that a memset() also isn't guaranteed to atomically update an unsigned long so it's theoretically possible to see torn and therefore bogus counter values. Link: https://lore.kernel.org/r/20241007-brauner-file-rcuref-v2-3-387e24dc9163@kernel.org Signed-off-by: Christian Brauner <brauner@kernel.org>
2024-10-07 14:23:59 +00:00
f = kmem_cache_alloc(filp_cachep, GFP_KERNEL);
if (unlikely(!f))
return ERR_PTR(-ENOMEM);
error = init_file(f, flags, cred);
if (unlikely(error)) {
kmem_cache_free(filp_cachep, f);
return ERR_PTR(error);
}
f->f_mode |= FMODE_NOACCOUNT;
return f;
}
/*
* Variant of alloc_empty_file() that allocates a backing_file container
* and doesn't check and modify nr_files.
*
* This is only for kernel internal use, and the allocate file must not be
* installed into file tables or such.
*/
struct file *alloc_empty_backing_file(int flags, const struct cred *cred)
{
struct backing_file *ff;
int error;
fs: port files to file_ref Port files to rely on file_ref reference to improve scaling and gain overflow protection. - We continue to WARN during get_file() in case a file that is already marked dead is revived as get_file() is only valid if the caller already holds a reference to the file. This hasn't changed just the check changes. - The semantics for epoll and ttm's dmabuf usage have changed. Both epoll and ttm synchronize with __fput() to prevent the underlying file from beeing freed. (1) epoll Explaining epoll is straightforward using a simple diagram. Essentially, the mutex of the epoll instance needs to be taken in both __fput() and around epi_fget() preventing the file from being freed while it is polled or preventing the file from being resurrected. CPU1 CPU2 fput(file) -> __fput(file) -> eventpoll_release(file) -> eventpoll_release_file(file) mutex_lock(&ep->mtx) epi_item_poll() -> epi_fget() -> file_ref_get(file) mutex_unlock(&ep->mtx) mutex_lock(&ep->mtx); __ep_remove() mutex_unlock(&ep->mtx); -> kmem_cache_free(file) (2) ttm dmabuf This explanation is a bit more involved. A regular dmabuf file stashed the dmabuf in file->private_data and the file in dmabuf->file: file->private_data = dmabuf; dmabuf->file = file; The generic release method of a dmabuf file handles file specific things: f_op->release::dma_buf_file_release() while the generic dentry release method of a dmabuf handles dmabuf freeing including driver specific things: dentry->d_release::dma_buf_release() During ttm dmabuf initialization in ttm_object_device_init() the ttm driver copies the provided struct dma_buf_ops into a private location: struct ttm_object_device { spinlock_t object_lock; struct dma_buf_ops ops; void (*dmabuf_release)(struct dma_buf *dma_buf); struct idr idr; }; ttm_object_device_init(const struct dma_buf_ops *ops) { // copy original dma_buf_ops in private location tdev->ops = *ops; // stash the release method of the original struct dma_buf_ops tdev->dmabuf_release = tdev->ops.release; // override the release method in the copy of the struct dma_buf_ops // with ttm's own dmabuf release method tdev->ops.release = ttm_prime_dmabuf_release; } When a new dmabuf is created the struct dma_buf_ops with the overriden release method set to ttm_prime_dmabuf_release is passed in exp_info.ops: DEFINE_DMA_BUF_EXPORT_INFO(exp_info); exp_info.ops = &tdev->ops; exp_info.size = prime->size; exp_info.flags = flags; exp_info.priv = prime; The call to dma_buf_export() then sets mutex_lock_interruptible(&prime->mutex); dma_buf = dma_buf_export(&exp_info) { dmabuf->ops = exp_info->ops; } mutex_unlock(&prime->mutex); which creates a new dmabuf file and then install a file descriptor to it in the callers file descriptor table: ret = dma_buf_fd(dma_buf, flags); When that dmabuf file is closed we now get: fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); Where we can see that prime->dma_buf is set to NULL. So when we have the following diagram: CPU1 CPU2 fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() ttm_prime_handle_to_fd() mutex_lock_interruptible(&prime->mutex) dma_buf = prime->dma_buf dma_buf && get_dma_buf_unless_doomed(dma_buf) -> file_ref_get(dma_buf->file) mutex_unlock(&prime->mutex); mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); -> kmem_cache_free(file) The logic of the mechanism is the same as for epoll: sync with __fput() preventing the file from being freed. Here the synchronization happens through the ttm instance's prime->mutex. Basically, the lifetime of the dma_buf and the file are tighly coupled. Both (1) and (2) used to call atomic_inc_not_zero() to check whether the file has already been marked dead and then refuse to revive it. This is only safe because both (1) and (2) sync with __fput() and thus prevent kmem_cache_free() on the file being called and thus prevent the file from being immediately recycled due to SLAB_TYPESAFE_BY_RCU. Both (1) and (2) have been ported from atomic_inc_not_zero() to file_ref_get(). That means a file that is already in the process of being marked as FILE_REF_DEAD: file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) can be revived again: CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) This is fine and inherent to the file_ref_get()/file_ref_put() semantics. For both (1) and (2) this is safe because __fput() is prevented from making progress if file_ref_get() fails due to the aforementioned synchronization mechanisms. Two cases need to be considered that affect both (1) epoll and (2) ttm dmabuf: (i) fput()'s file_ref_put() and marks the file as FILE_REF_NOREF but before that fput() can mark the file as FILE_REF_DEAD someone manages to sneak in a file_ref_get() and brings the refcount back from FILE_REF_NOREF to FILE_REF_ONEREF. In that case the original fput() doesn't call __fput(). For epoll the poll will finish and for ttm dmabuf the file can be used again. For ttm dambuf this is actually an advantage because it avoids immediately allocating a new dmabuf object. CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) (ii) fput()'s file_ref_put() marks the file FILE_REF_NOREF and also suceeds in actually marking it FILE_REF_DEAD and then calls into __fput() to free the file. When either (1) or (2) call file_ref_get() they fail as atomic_long_add_negative() will return true. At the same time, both (1) and (2) all file_ref_get() under mutexes that __fput() must also acquire preventing kmem_cache_free() from freeing the file. So while this might be treated as a change in semantics for (1) and (2) it really isn't. It if should end up causing issues this can be fixed by adding a helper that does something like: long cnt = atomic_long_read(&ref->refcnt); do { if (cnt < 0) return false; } while (!atomic_long_try_cmpxchg(&ref->refcnt, &cnt, cnt + 1)); return true; which would block FILE_REF_NOREF to FILE_REF_ONEREF transitions. - Jann correctly pointed out that kmem_cache_zalloc() cannot be used anymore once files have been ported to file_ref_t. The kmem_cache_zalloc() call will memset() the whole struct file to zero when it is reallocated. This will also set file->f_ref to zero which mens that a concurrent file_ref_get() can return true: CPU1 CPU2 __get_file_rcu() rcu_dereference_raw() close() [frees file] alloc_empty_file() kmem_cache_zalloc() [reallocates same file] memset(..., 0, ...) file_ref_get() [increments 0->1, returns true] init_file() file_ref_init(..., 1) [sets to 0] rcu_dereference_raw() fput() file_ref_put() [decrements 0->FILE_REF_NOREF, frees file] [UAF] causing a concurrent __get_file_rcu() call to acquire a reference to the file that is about to be reallocated and immediately freeing it on realizing that it has been recycled. This causes a UAF for the task that reallocated/recycled the file. This is prevented by switching from kmem_cache_zalloc() to kmem_cache_alloc() and initializing the fields manually. With file->f_ref initialized last. Note that a memset() also isn't guaranteed to atomically update an unsigned long so it's theoretically possible to see torn and therefore bogus counter values. Link: https://lore.kernel.org/r/20241007-brauner-file-rcuref-v2-3-387e24dc9163@kernel.org Signed-off-by: Christian Brauner <brauner@kernel.org>
2024-10-07 14:23:59 +00:00
ff = kmem_cache_alloc(bfilp_cachep, GFP_KERNEL);
if (unlikely(!ff))
return ERR_PTR(-ENOMEM);
error = init_file(&ff->file, flags, cred);
if (unlikely(error)) {
kmem_cache_free(bfilp_cachep, ff);
return ERR_PTR(error);
}
ff->file.f_mode |= FMODE_BACKING | FMODE_NOACCOUNT;
return &ff->file;
}
r/o bind mounts: filesystem helpers for custom 'struct file's Why do we need r/o bind mounts? This feature allows a read-only view into a read-write filesystem. In the process of doing that, it also provides infrastructure for keeping track of the number of writers to any given mount. This has a number of uses. It allows chroots to have parts of filesystems writable. It will be useful for containers in the future because users may have root inside a container, but should not be allowed to write to somefilesystems. This also replaces patches that vserver has had out of the tree for several years. It allows security enhancement by making sure that parts of your filesystem read-only (such as when you don't trust your FTP server), when you don't want to have entire new filesystems mounted, or when you want atime selectively updated. I've been using the following script to test that the feature is working as desired. It takes a directory and makes a regular bind and a r/o bind mount of it. It then performs some normal filesystem operations on the three directories, including ones that are expected to fail, like creating a file on the r/o mount. This patch: Some filesystems forego the vfs and may_open() and create their own 'struct file's. This patch creates a couple of helper functions which can be used by these filesystems, and will provide a unified place which the r/o bind mount code may patch. Also, rename an existing, static-scope init_file() to a less generic name. Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Cc: Christoph Hellwig <hch@lst.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-17 06:31:13 +00:00
/**
* file_init_path - initialize a 'struct file' based on path
*
* @file: the file to set up
* @path: the (dentry, vfsmount) pair for the new file
r/o bind mounts: filesystem helpers for custom 'struct file's Why do we need r/o bind mounts? This feature allows a read-only view into a read-write filesystem. In the process of doing that, it also provides infrastructure for keeping track of the number of writers to any given mount. This has a number of uses. It allows chroots to have parts of filesystems writable. It will be useful for containers in the future because users may have root inside a container, but should not be allowed to write to somefilesystems. This also replaces patches that vserver has had out of the tree for several years. It allows security enhancement by making sure that parts of your filesystem read-only (such as when you don't trust your FTP server), when you don't want to have entire new filesystems mounted, or when you want atime selectively updated. I've been using the following script to test that the feature is working as desired. It takes a directory and makes a regular bind and a r/o bind mount of it. It then performs some normal filesystem operations on the three directories, including ones that are expected to fail, like creating a file on the r/o mount. This patch: Some filesystems forego the vfs and may_open() and create their own 'struct file's. This patch creates a couple of helper functions which can be used by these filesystems, and will provide a unified place which the r/o bind mount code may patch. Also, rename an existing, static-scope init_file() to a less generic name. Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Cc: Christoph Hellwig <hch@lst.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-17 06:31:13 +00:00
* @fop: the 'struct file_operations' for the new file
*/
static void file_init_path(struct file *file, const struct path *path,
const struct file_operations *fop)
r/o bind mounts: filesystem helpers for custom 'struct file's Why do we need r/o bind mounts? This feature allows a read-only view into a read-write filesystem. In the process of doing that, it also provides infrastructure for keeping track of the number of writers to any given mount. This has a number of uses. It allows chroots to have parts of filesystems writable. It will be useful for containers in the future because users may have root inside a container, but should not be allowed to write to somefilesystems. This also replaces patches that vserver has had out of the tree for several years. It allows security enhancement by making sure that parts of your filesystem read-only (such as when you don't trust your FTP server), when you don't want to have entire new filesystems mounted, or when you want atime selectively updated. I've been using the following script to test that the feature is working as desired. It takes a directory and makes a regular bind and a r/o bind mount of it. It then performs some normal filesystem operations on the three directories, including ones that are expected to fail, like creating a file on the r/o mount. This patch: Some filesystems forego the vfs and may_open() and create their own 'struct file's. This patch creates a couple of helper functions which can be used by these filesystems, and will provide a unified place which the r/o bind mount code may patch. Also, rename an existing, static-scope init_file() to a less generic name. Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Cc: Christoph Hellwig <hch@lst.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-17 06:31:13 +00:00
{
file->f_path = *path;
file->f_inode = path->dentry->d_inode;
file->f_mapping = path->dentry->d_inode->i_mapping;
fs: new infrastructure for writeback error handling and reporting Most filesystems currently use mapping_set_error and filemap_check_errors for setting and reporting/clearing writeback errors at the mapping level. filemap_check_errors is indirectly called from most of the filemap_fdatawait_* functions and from filemap_write_and_wait*. These functions are called from all sorts of contexts to wait on writeback to finish -- e.g. mostly in fsync, but also in truncate calls, getattr, etc. The non-fsync callers are problematic. We should be reporting writeback errors during fsync, but many places spread over the tree clear out errors before they can be properly reported, or report errors at nonsensical times. If I get -EIO on a stat() call, there is no reason for me to assume that it is because some previous writeback failed. The fact that it also clears out the error such that a subsequent fsync returns 0 is a bug, and a nasty one since that's potentially silent data corruption. This patch adds a small bit of new infrastructure for setting and reporting errors during address_space writeback. While the above was my original impetus for adding this, I think it's also the case that current fsync semantics are just problematic for userland. Most applications that call fsync do so to ensure that the data they wrote has hit the backing store. In the case where there are multiple writers to the file at the same time, this is really hard to determine. The first one to call fsync will see any stored error, and the rest get back 0. The processes with open fds may not be associated with one another in any way. They could even be in different containers, so ensuring coordination between all fsync callers is not really an option. One way to remedy this would be to track what file descriptor was used to dirty the file, but that's rather cumbersome and would likely be slow. However, there is a simpler way to improve the semantics here without incurring too much overhead. This set adds an errseq_t to struct address_space, and a corresponding one is added to struct file. Writeback errors are recorded in the mapping's errseq_t, and the one in struct file is used as the "since" value. This changes the semantics of the Linux fsync implementation such that applications can now use it to determine whether there were any writeback errors since fsync(fd) was last called (or since the file was opened in the case of fsync having never been called). Note that those writeback errors may have occurred when writing data that was dirtied via an entirely different fd, but that's the case now with the current mapping_set_error/filemap_check_error infrastructure. This will at least prevent you from getting a false report of success. The new behavior is still consistent with the POSIX spec, and is more reliable for application developers. This patch just adds some basic infrastructure for doing this, and ensures that the f_wb_err "cursor" is properly set when a file is opened. Later patches will change the existing code to use this new infrastructure for reporting errors at fsync time. Signed-off-by: Jeff Layton <jlayton@redhat.com> Reviewed-by: Jan Kara <jack@suse.cz>
2017-07-06 11:02:25 +00:00
file->f_wb_err = filemap_sample_wb_err(file->f_mapping);
vfs: track per-sb writeback errors and report them to syncfs Patch series "vfs: have syncfs() return error when there are writeback errors", v6. Currently, syncfs does not return errors when one of the inodes fails to be written back. It will return errors based on the legacy AS_EIO and AS_ENOSPC flags when syncing out the block device fails, but that's not particularly helpful for filesystems that aren't backed by a blockdev. It's also possible for a stray sync to lose those errors. The basic idea in this set is to track writeback errors at the superblock level, so that we can quickly and easily check whether something bad happened without having to fsync each file individually. syncfs is then changed to reliably report writeback errors after they occur, much in the same fashion as fsync does now. This patch (of 2): Usually we suggest that applications call fsync when they want to ensure that all data written to the file has made it to the backing store, but that can be inefficient when there are a lot of open files. Calling syncfs on the filesystem can be more efficient in some situations, but the error reporting doesn't currently work the way most people expect. If a single inode on a filesystem reports a writeback error, syncfs won't necessarily return an error. syncfs only returns an error if __sync_blockdev fails, and on some filesystems that's a no-op. It would be better if syncfs reported an error if there were any writeback failures. Then applications could call syncfs to see if there are any errors on any open files, and could then call fsync on all of the other descriptors to figure out which one failed. This patch adds a new errseq_t to struct super_block, and has mapping_set_error also record writeback errors there. To report those errors, we also need to keep an errseq_t in struct file to act as a cursor. This patch adds a dedicated field for that purpose, which slots nicely into 4 bytes of padding at the end of struct file on x86_64. An earlier version of this patch used an O_PATH file descriptor to cue the kernel that the open file should track the superblock error and not the inode's writeback error. I think that API is just too weird though. This is simpler and should make syncfs error reporting "just work" even if someone is multiplexing fsync and syncfs on the same fds. Signed-off-by: Jeff Layton <jlayton@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Reviewed-by: Jan Kara <jack@suse.cz> Cc: Andres Freund <andres@anarazel.de> Cc: Matthew Wilcox <willy@infradead.org> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Christoph Hellwig <hch@infradead.org> Cc: Dave Chinner <david@fromorbit.com> Cc: David Howells <dhowells@redhat.com> Link: http://lkml.kernel.org/r/20200428135155.19223-1-jlayton@kernel.org Link: http://lkml.kernel.org/r/20200428135155.19223-2-jlayton@kernel.org Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2020-06-02 04:45:36 +00:00
file->f_sb_err = file_sample_sb_err(file);
if (fop->llseek)
file->f_mode |= FMODE_LSEEK;
if ((file->f_mode & FMODE_READ) &&
likely(fop->read || fop->read_iter))
file->f_mode |= FMODE_CAN_READ;
if ((file->f_mode & FMODE_WRITE) &&
likely(fop->write || fop->write_iter))
file->f_mode |= FMODE_CAN_WRITE;
file->f_iocb_flags = iocb_flags(file);
file->f_mode |= FMODE_OPENED;
r/o bind mounts: filesystem helpers for custom 'struct file's Why do we need r/o bind mounts? This feature allows a read-only view into a read-write filesystem. In the process of doing that, it also provides infrastructure for keeping track of the number of writers to any given mount. This has a number of uses. It allows chroots to have parts of filesystems writable. It will be useful for containers in the future because users may have root inside a container, but should not be allowed to write to somefilesystems. This also replaces patches that vserver has had out of the tree for several years. It allows security enhancement by making sure that parts of your filesystem read-only (such as when you don't trust your FTP server), when you don't want to have entire new filesystems mounted, or when you want atime selectively updated. I've been using the following script to test that the feature is working as desired. It takes a directory and makes a regular bind and a r/o bind mount of it. It then performs some normal filesystem operations on the three directories, including ones that are expected to fail, like creating a file on the r/o mount. This patch: Some filesystems forego the vfs and may_open() and create their own 'struct file's. This patch creates a couple of helper functions which can be used by these filesystems, and will provide a unified place which the r/o bind mount code may patch. Also, rename an existing, static-scope init_file() to a less generic name. Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Cc: Christoph Hellwig <hch@lst.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-17 06:31:13 +00:00
file->f_op = fop;
if ((file->f_mode & (FMODE_READ | FMODE_WRITE)) == FMODE_READ)
i_readcount_inc(path->dentry->d_inode);
}
/**
* alloc_file - allocate and initialize a 'struct file'
*
* @path: the (dentry, vfsmount) pair for the new file
* @flags: O_... flags with which the new file will be opened
* @fop: the 'struct file_operations' for the new file
*/
static struct file *alloc_file(const struct path *path, int flags,
const struct file_operations *fop)
{
struct file *file;
file = alloc_empty_file(flags, current_cred());
if (!IS_ERR(file))
file_init_path(file, path, fop);
return file;
r/o bind mounts: filesystem helpers for custom 'struct file's Why do we need r/o bind mounts? This feature allows a read-only view into a read-write filesystem. In the process of doing that, it also provides infrastructure for keeping track of the number of writers to any given mount. This has a number of uses. It allows chroots to have parts of filesystems writable. It will be useful for containers in the future because users may have root inside a container, but should not be allowed to write to somefilesystems. This also replaces patches that vserver has had out of the tree for several years. It allows security enhancement by making sure that parts of your filesystem read-only (such as when you don't trust your FTP server), when you don't want to have entire new filesystems mounted, or when you want atime selectively updated. I've been using the following script to test that the feature is working as desired. It takes a directory and makes a regular bind and a r/o bind mount of it. It then performs some normal filesystem operations on the three directories, including ones that are expected to fail, like creating a file on the r/o mount. This patch: Some filesystems forego the vfs and may_open() and create their own 'struct file's. This patch creates a couple of helper functions which can be used by these filesystems, and will provide a unified place which the r/o bind mount code may patch. Also, rename an existing, static-scope init_file() to a less generic name. Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Cc: Christoph Hellwig <hch@lst.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-17 06:31:13 +00:00
}
static inline int alloc_path_pseudo(const char *name, struct inode *inode,
struct vfsmount *mnt, struct path *path)
{
struct qstr this = QSTR_INIT(name, strlen(name));
path->dentry = d_alloc_pseudo(mnt->mnt_sb, &this);
if (!path->dentry)
return -ENOMEM;
path->mnt = mntget(mnt);
d_instantiate(path->dentry, inode);
return 0;
}
struct file *alloc_file_pseudo(struct inode *inode, struct vfsmount *mnt,
const char *name, int flags,
const struct file_operations *fops)
{
int ret;
struct path path;
struct file *file;
ret = alloc_path_pseudo(name, inode, mnt, &path);
if (ret)
return ERR_PTR(ret);
2020-06-29 14:41:45 +00:00
file = alloc_file(&path, flags, fops);
if (IS_ERR(file)) {
ihold(inode);
path_put(&path);
}
return file;
}
EXPORT_SYMBOL(alloc_file_pseudo);
struct file *alloc_file_pseudo_noaccount(struct inode *inode,
struct vfsmount *mnt, const char *name,
int flags,
const struct file_operations *fops)
{
int ret;
struct path path;
struct file *file;
ret = alloc_path_pseudo(name, inode, mnt, &path);
if (ret)
return ERR_PTR(ret);
file = alloc_empty_file_noaccount(flags, current_cred());
if (IS_ERR(file)) {
ihold(inode);
path_put(&path);
return file;
}
file_init_path(file, &path, fops);
return file;
}
EXPORT_SYMBOL_GPL(alloc_file_pseudo_noaccount);
struct file *alloc_file_clone(struct file *base, int flags,
const struct file_operations *fops)
{
struct file *f;
f = alloc_file(&base->f_path, flags, fops);
if (!IS_ERR(f)) {
path_get(&f->f_path);
f->f_mapping = base->f_mapping;
}
return f;
}
/* the real guts of fput() - releasing the last reference to file
*/
static void __fput(struct file *file)
{
struct dentry *dentry = file->f_path.dentry;
struct vfsmount *mnt = file->f_path.mnt;
struct inode *inode = file->f_inode;
fmode_t mode = file->f_mode;
if (unlikely(!(file->f_mode & FMODE_OPENED)))
goto out;
might_sleep();
fsnotify_close(file);
/*
* The function eventpoll_release() should be the first called
* in the file cleanup chain.
*/
eventpoll_release(file);
locks_remove_file(file);
security_file_release(file);
if (unlikely(file->f_flags & FASYNC)) {
if (file->f_op->fasync)
file->f_op->fasync(-1, file, 0);
}
if (file->f_op->release)
file->f_op->release(inode, file);
if (unlikely(S_ISCHR(inode->i_mode) && inode->i_cdev != NULL &&
!(mode & FMODE_PATH))) {
cdev_put(inode->i_cdev);
}
fops_put(file->f_op);
file: reclaim 24 bytes from f_owner We do embedd struct fown_struct into struct file letting it take up 32 bytes in total. We could tweak struct fown_struct to be more compact but really it shouldn't even be embedded in struct file in the first place. Instead, actual users of struct fown_struct should allocate the struct on demand. This frees up 24 bytes in struct file. That will have some potentially user-visible changes for the ownership fcntl()s. Some of them can now fail due to allocation failures. Practically, that probably will almost never happen as the allocations are small and they only happen once per file. The fown_struct is used during kill_fasync() which is used by e.g., pipes to generate a SIGIO signal. Sending of such signals is conditional on userspace having set an owner for the file using one of the F_OWNER fcntl()s. Such users will be unaffected if struct fown_struct is allocated during the fcntl() call. There are a few subsystems that call __f_setown() expecting file->f_owner to be allocated: (1) tun devices file->f_op->fasync::tun_chr_fasync() -> __f_setown() There are no callers of tun_chr_fasync(). (2) tty devices file->f_op->fasync::tty_fasync() -> __tty_fasync() -> __f_setown() tty_fasync() has no additional callers but __tty_fasync() has. Note that __tty_fasync() only calls __f_setown() if the @on argument is true. It's called from: file->f_op->release::tty_release() -> tty_release() -> __tty_fasync() -> __f_setown() tty_release() calls __tty_fasync() with @on false => __f_setown() is never called from tty_release(). => All callers of tty_release() are safe as well. file->f_op->release::tty_open() -> tty_release() -> __tty_fasync() -> __f_setown() __tty_hangup() calls __tty_fasync() with @on false => __f_setown() is never called from tty_release(). => All callers of __tty_hangup() are safe as well. From the callchains it's obvious that (1) and (2) end up getting called via file->f_op->fasync(). That can happen either through the F_SETFL fcntl() with the FASYNC flag raised or via the FIOASYNC ioctl(). If FASYNC is requested and the file isn't already FASYNC then file->f_op->fasync() is called with @on true which ends up causing both (1) and (2) to call __f_setown(). (1) and (2) are the only subsystems that call __f_setown() from the file->f_op->fasync() handler. So both (1) and (2) have been updated to allocate a struct fown_struct prior to calling fasync_helper() to register with the fasync infrastructure. That's safe as they both call fasync_helper() which also does allocations if @on is true. The other interesting case are file leases: (3) file leases lease_manager_ops->lm_setup::lease_setup() -> __f_setown() Which in turn is called from: generic_add_lease() -> lease_manager_ops->lm_setup::lease_setup() -> __f_setown() So here again we can simply make generic_add_lease() allocate struct fown_struct prior to the lease_manager_ops->lm_setup::lease_setup() which happens under a spinlock. With that the two remaining subsystems that call __f_setown() are: (4) dnotify (5) sockets Both have their own custom ioctls to set struct fown_struct and both have been converted to allocate a struct fown_struct on demand from their respective ioctls. Interactions with O_PATH are fine as well e.g., when opening a /dev/tty as O_PATH then no file->f_op->open() happens thus no file->f_owner is allocated. That's fine as no file operation will be set for those and the device has never been opened. fcntl()s called on such things will just allocate a ->f_owner on demand. Although I have zero idea why'd you care about f_owner on an O_PATH fd. Link: https://lore.kernel.org/r/20240813-work-f_owner-v2-1-4e9343a79f9f@kernel.org Reviewed-by: Jeff Layton <jlayton@kernel.org> Signed-off-by: Christian Brauner <brauner@kernel.org>
2024-08-09 16:00:01 +00:00
file_f_owner_release(file);
put_file_access(file);
dput(dentry);
if (unlikely(mode & FMODE_NEED_UNMOUNT))
dissolve_on_fput(mnt);
mntput(mnt);
out:
file_free(file);
}
static LLIST_HEAD(delayed_fput_list);
static void delayed_fput(struct work_struct *unused)
{
struct llist_node *node = llist_del_all(&delayed_fput_list);
struct file *f, *t;
llist_for_each_entry_safe(f, t, node, f_llist)
__fput(f);
}
static void ____fput(struct callback_head *work)
{
__fput(container_of(work, struct file, f_task_work));
}
/*
* If kernel thread really needs to have the final fput() it has done
* to complete, call this. The only user right now is the boot - we
* *do* need to make sure our writes to binaries on initramfs has
* not left us with opened struct file waiting for __fput() - execve()
* won't work without that. Please, don't add more callers without
* very good reasons; in particular, never call that with locks
* held and never call that from a thread that might need to do
* some work on any kind of umount.
*/
void flush_delayed_fput(void)
{
delayed_fput(NULL);
}
EXPORT_SYMBOL_GPL(flush_delayed_fput);
static DECLARE_DELAYED_WORK(delayed_fput_work, delayed_fput);
void fput(struct file *file)
{
fs: port files to file_ref Port files to rely on file_ref reference to improve scaling and gain overflow protection. - We continue to WARN during get_file() in case a file that is already marked dead is revived as get_file() is only valid if the caller already holds a reference to the file. This hasn't changed just the check changes. - The semantics for epoll and ttm's dmabuf usage have changed. Both epoll and ttm synchronize with __fput() to prevent the underlying file from beeing freed. (1) epoll Explaining epoll is straightforward using a simple diagram. Essentially, the mutex of the epoll instance needs to be taken in both __fput() and around epi_fget() preventing the file from being freed while it is polled or preventing the file from being resurrected. CPU1 CPU2 fput(file) -> __fput(file) -> eventpoll_release(file) -> eventpoll_release_file(file) mutex_lock(&ep->mtx) epi_item_poll() -> epi_fget() -> file_ref_get(file) mutex_unlock(&ep->mtx) mutex_lock(&ep->mtx); __ep_remove() mutex_unlock(&ep->mtx); -> kmem_cache_free(file) (2) ttm dmabuf This explanation is a bit more involved. A regular dmabuf file stashed the dmabuf in file->private_data and the file in dmabuf->file: file->private_data = dmabuf; dmabuf->file = file; The generic release method of a dmabuf file handles file specific things: f_op->release::dma_buf_file_release() while the generic dentry release method of a dmabuf handles dmabuf freeing including driver specific things: dentry->d_release::dma_buf_release() During ttm dmabuf initialization in ttm_object_device_init() the ttm driver copies the provided struct dma_buf_ops into a private location: struct ttm_object_device { spinlock_t object_lock; struct dma_buf_ops ops; void (*dmabuf_release)(struct dma_buf *dma_buf); struct idr idr; }; ttm_object_device_init(const struct dma_buf_ops *ops) { // copy original dma_buf_ops in private location tdev->ops = *ops; // stash the release method of the original struct dma_buf_ops tdev->dmabuf_release = tdev->ops.release; // override the release method in the copy of the struct dma_buf_ops // with ttm's own dmabuf release method tdev->ops.release = ttm_prime_dmabuf_release; } When a new dmabuf is created the struct dma_buf_ops with the overriden release method set to ttm_prime_dmabuf_release is passed in exp_info.ops: DEFINE_DMA_BUF_EXPORT_INFO(exp_info); exp_info.ops = &tdev->ops; exp_info.size = prime->size; exp_info.flags = flags; exp_info.priv = prime; The call to dma_buf_export() then sets mutex_lock_interruptible(&prime->mutex); dma_buf = dma_buf_export(&exp_info) { dmabuf->ops = exp_info->ops; } mutex_unlock(&prime->mutex); which creates a new dmabuf file and then install a file descriptor to it in the callers file descriptor table: ret = dma_buf_fd(dma_buf, flags); When that dmabuf file is closed we now get: fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); Where we can see that prime->dma_buf is set to NULL. So when we have the following diagram: CPU1 CPU2 fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() ttm_prime_handle_to_fd() mutex_lock_interruptible(&prime->mutex) dma_buf = prime->dma_buf dma_buf && get_dma_buf_unless_doomed(dma_buf) -> file_ref_get(dma_buf->file) mutex_unlock(&prime->mutex); mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); -> kmem_cache_free(file) The logic of the mechanism is the same as for epoll: sync with __fput() preventing the file from being freed. Here the synchronization happens through the ttm instance's prime->mutex. Basically, the lifetime of the dma_buf and the file are tighly coupled. Both (1) and (2) used to call atomic_inc_not_zero() to check whether the file has already been marked dead and then refuse to revive it. This is only safe because both (1) and (2) sync with __fput() and thus prevent kmem_cache_free() on the file being called and thus prevent the file from being immediately recycled due to SLAB_TYPESAFE_BY_RCU. Both (1) and (2) have been ported from atomic_inc_not_zero() to file_ref_get(). That means a file that is already in the process of being marked as FILE_REF_DEAD: file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) can be revived again: CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) This is fine and inherent to the file_ref_get()/file_ref_put() semantics. For both (1) and (2) this is safe because __fput() is prevented from making progress if file_ref_get() fails due to the aforementioned synchronization mechanisms. Two cases need to be considered that affect both (1) epoll and (2) ttm dmabuf: (i) fput()'s file_ref_put() and marks the file as FILE_REF_NOREF but before that fput() can mark the file as FILE_REF_DEAD someone manages to sneak in a file_ref_get() and brings the refcount back from FILE_REF_NOREF to FILE_REF_ONEREF. In that case the original fput() doesn't call __fput(). For epoll the poll will finish and for ttm dmabuf the file can be used again. For ttm dambuf this is actually an advantage because it avoids immediately allocating a new dmabuf object. CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) (ii) fput()'s file_ref_put() marks the file FILE_REF_NOREF and also suceeds in actually marking it FILE_REF_DEAD and then calls into __fput() to free the file. When either (1) or (2) call file_ref_get() they fail as atomic_long_add_negative() will return true. At the same time, both (1) and (2) all file_ref_get() under mutexes that __fput() must also acquire preventing kmem_cache_free() from freeing the file. So while this might be treated as a change in semantics for (1) and (2) it really isn't. It if should end up causing issues this can be fixed by adding a helper that does something like: long cnt = atomic_long_read(&ref->refcnt); do { if (cnt < 0) return false; } while (!atomic_long_try_cmpxchg(&ref->refcnt, &cnt, cnt + 1)); return true; which would block FILE_REF_NOREF to FILE_REF_ONEREF transitions. - Jann correctly pointed out that kmem_cache_zalloc() cannot be used anymore once files have been ported to file_ref_t. The kmem_cache_zalloc() call will memset() the whole struct file to zero when it is reallocated. This will also set file->f_ref to zero which mens that a concurrent file_ref_get() can return true: CPU1 CPU2 __get_file_rcu() rcu_dereference_raw() close() [frees file] alloc_empty_file() kmem_cache_zalloc() [reallocates same file] memset(..., 0, ...) file_ref_get() [increments 0->1, returns true] init_file() file_ref_init(..., 1) [sets to 0] rcu_dereference_raw() fput() file_ref_put() [decrements 0->FILE_REF_NOREF, frees file] [UAF] causing a concurrent __get_file_rcu() call to acquire a reference to the file that is about to be reallocated and immediately freeing it on realizing that it has been recycled. This causes a UAF for the task that reallocated/recycled the file. This is prevented by switching from kmem_cache_zalloc() to kmem_cache_alloc() and initializing the fields manually. With file->f_ref initialized last. Note that a memset() also isn't guaranteed to atomically update an unsigned long so it's theoretically possible to see torn and therefore bogus counter values. Link: https://lore.kernel.org/r/20241007-brauner-file-rcuref-v2-3-387e24dc9163@kernel.org Signed-off-by: Christian Brauner <brauner@kernel.org>
2024-10-07 14:23:59 +00:00
if (file_ref_put(&file->f_ref)) {
struct task_struct *task = current;
if (unlikely(!(file->f_mode & (FMODE_BACKING | FMODE_OPENED)))) {
file_free(file);
return;
}
if (likely(!in_interrupt() && !(task->flags & PF_KTHREAD))) {
init_task_work(&file->f_task_work, ____fput);
if (!task_work_add(task, &file->f_task_work, TWA_RESUME))
return;
/*
* After this task has run exit_task_work(),
* task_work_add() will fail. Fall through to delayed
* fput to avoid leaking *file.
*/
}
if (llist_add(&file->f_llist, &delayed_fput_list))
schedule_delayed_work(&delayed_fput_work, 1);
}
}
/*
* synchronous analog of fput(); for kernel threads that might be needed
* in some umount() (and thus can't use flush_delayed_fput() without
* risking deadlocks), need to wait for completion of __fput() and know
* for this specific struct file it won't involve anything that would
* need them. Use only if you really need it - at the very least,
* don't blindly convert fput() by kernel thread to that.
*/
void __fput_sync(struct file *file)
{
fs: port files to file_ref Port files to rely on file_ref reference to improve scaling and gain overflow protection. - We continue to WARN during get_file() in case a file that is already marked dead is revived as get_file() is only valid if the caller already holds a reference to the file. This hasn't changed just the check changes. - The semantics for epoll and ttm's dmabuf usage have changed. Both epoll and ttm synchronize with __fput() to prevent the underlying file from beeing freed. (1) epoll Explaining epoll is straightforward using a simple diagram. Essentially, the mutex of the epoll instance needs to be taken in both __fput() and around epi_fget() preventing the file from being freed while it is polled or preventing the file from being resurrected. CPU1 CPU2 fput(file) -> __fput(file) -> eventpoll_release(file) -> eventpoll_release_file(file) mutex_lock(&ep->mtx) epi_item_poll() -> epi_fget() -> file_ref_get(file) mutex_unlock(&ep->mtx) mutex_lock(&ep->mtx); __ep_remove() mutex_unlock(&ep->mtx); -> kmem_cache_free(file) (2) ttm dmabuf This explanation is a bit more involved. A regular dmabuf file stashed the dmabuf in file->private_data and the file in dmabuf->file: file->private_data = dmabuf; dmabuf->file = file; The generic release method of a dmabuf file handles file specific things: f_op->release::dma_buf_file_release() while the generic dentry release method of a dmabuf handles dmabuf freeing including driver specific things: dentry->d_release::dma_buf_release() During ttm dmabuf initialization in ttm_object_device_init() the ttm driver copies the provided struct dma_buf_ops into a private location: struct ttm_object_device { spinlock_t object_lock; struct dma_buf_ops ops; void (*dmabuf_release)(struct dma_buf *dma_buf); struct idr idr; }; ttm_object_device_init(const struct dma_buf_ops *ops) { // copy original dma_buf_ops in private location tdev->ops = *ops; // stash the release method of the original struct dma_buf_ops tdev->dmabuf_release = tdev->ops.release; // override the release method in the copy of the struct dma_buf_ops // with ttm's own dmabuf release method tdev->ops.release = ttm_prime_dmabuf_release; } When a new dmabuf is created the struct dma_buf_ops with the overriden release method set to ttm_prime_dmabuf_release is passed in exp_info.ops: DEFINE_DMA_BUF_EXPORT_INFO(exp_info); exp_info.ops = &tdev->ops; exp_info.size = prime->size; exp_info.flags = flags; exp_info.priv = prime; The call to dma_buf_export() then sets mutex_lock_interruptible(&prime->mutex); dma_buf = dma_buf_export(&exp_info) { dmabuf->ops = exp_info->ops; } mutex_unlock(&prime->mutex); which creates a new dmabuf file and then install a file descriptor to it in the callers file descriptor table: ret = dma_buf_fd(dma_buf, flags); When that dmabuf file is closed we now get: fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); Where we can see that prime->dma_buf is set to NULL. So when we have the following diagram: CPU1 CPU2 fput(file) -> __fput(file) -> f_op->release::dma_buf_file_release() -> dput() -> d_op->d_release::dma_buf_release() -> dmabuf->ops->release::ttm_prime_dmabuf_release() ttm_prime_handle_to_fd() mutex_lock_interruptible(&prime->mutex) dma_buf = prime->dma_buf dma_buf && get_dma_buf_unless_doomed(dma_buf) -> file_ref_get(dma_buf->file) mutex_unlock(&prime->mutex); mutex_lock(&prime->mutex); if (prime->dma_buf == dma_buf) prime->dma_buf = NULL; mutex_unlock(&prime->mutex); -> kmem_cache_free(file) The logic of the mechanism is the same as for epoll: sync with __fput() preventing the file from being freed. Here the synchronization happens through the ttm instance's prime->mutex. Basically, the lifetime of the dma_buf and the file are tighly coupled. Both (1) and (2) used to call atomic_inc_not_zero() to check whether the file has already been marked dead and then refuse to revive it. This is only safe because both (1) and (2) sync with __fput() and thus prevent kmem_cache_free() on the file being called and thus prevent the file from being immediately recycled due to SLAB_TYPESAFE_BY_RCU. Both (1) and (2) have been ported from atomic_inc_not_zero() to file_ref_get(). That means a file that is already in the process of being marked as FILE_REF_DEAD: file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) can be revived again: CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) This is fine and inherent to the file_ref_get()/file_ref_put() semantics. For both (1) and (2) this is safe because __fput() is prevented from making progress if file_ref_get() fails due to the aforementioned synchronization mechanisms. Two cases need to be considered that affect both (1) epoll and (2) ttm dmabuf: (i) fput()'s file_ref_put() and marks the file as FILE_REF_NOREF but before that fput() can mark the file as FILE_REF_DEAD someone manages to sneak in a file_ref_get() and brings the refcount back from FILE_REF_NOREF to FILE_REF_ONEREF. In that case the original fput() doesn't call __fput(). For epoll the poll will finish and for ttm dmabuf the file can be used again. For ttm dambuf this is actually an advantage because it avoids immediately allocating a new dmabuf object. CPU1 CPU2 file_ref_put() cnt = atomic_long_dec_return() -> __file_ref_put(cnt) if (cnt == FIlE_REF_NOREF) file_ref_get() // Brings reference back to FILE_REF_ONEREF atomic_long_add_negative() atomic_long_try_cmpxchg_release(cnt, FILE_REF_DEAD) (ii) fput()'s file_ref_put() marks the file FILE_REF_NOREF and also suceeds in actually marking it FILE_REF_DEAD and then calls into __fput() to free the file. When either (1) or (2) call file_ref_get() they fail as atomic_long_add_negative() will return true. At the same time, both (1) and (2) all file_ref_get() under mutexes that __fput() must also acquire preventing kmem_cache_free() from freeing the file. So while this might be treated as a change in semantics for (1) and (2) it really isn't. It if should end up causing issues this can be fixed by adding a helper that does something like: long cnt = atomic_long_read(&ref->refcnt); do { if (cnt < 0) return false; } while (!atomic_long_try_cmpxchg(&ref->refcnt, &cnt, cnt + 1)); return true; which would block FILE_REF_NOREF to FILE_REF_ONEREF transitions. - Jann correctly pointed out that kmem_cache_zalloc() cannot be used anymore once files have been ported to file_ref_t. The kmem_cache_zalloc() call will memset() the whole struct file to zero when it is reallocated. This will also set file->f_ref to zero which mens that a concurrent file_ref_get() can return true: CPU1 CPU2 __get_file_rcu() rcu_dereference_raw() close() [frees file] alloc_empty_file() kmem_cache_zalloc() [reallocates same file] memset(..., 0, ...) file_ref_get() [increments 0->1, returns true] init_file() file_ref_init(..., 1) [sets to 0] rcu_dereference_raw() fput() file_ref_put() [decrements 0->FILE_REF_NOREF, frees file] [UAF] causing a concurrent __get_file_rcu() call to acquire a reference to the file that is about to be reallocated and immediately freeing it on realizing that it has been recycled. This causes a UAF for the task that reallocated/recycled the file. This is prevented by switching from kmem_cache_zalloc() to kmem_cache_alloc() and initializing the fields manually. With file->f_ref initialized last. Note that a memset() also isn't guaranteed to atomically update an unsigned long so it's theoretically possible to see torn and therefore bogus counter values. Link: https://lore.kernel.org/r/20241007-brauner-file-rcuref-v2-3-387e24dc9163@kernel.org Signed-off-by: Christian Brauner <brauner@kernel.org>
2024-10-07 14:23:59 +00:00
if (file_ref_put(&file->f_ref))
__fput(file);
}
EXPORT_SYMBOL(fput);
EXPORT_SYMBOL(__fput_sync);
void __init files_init(void)
{
struct kmem_cache_args args = {
.use_freeptr_offset = true,
.freeptr_offset = offsetof(struct file, f_freeptr),
};
filp_cachep = kmem_cache_create("filp", sizeof(struct file), &args,
SLAB_HWCACHE_ALIGN | SLAB_PANIC |
SLAB_ACCOUNT | SLAB_TYPESAFE_BY_RCU);
args.freeptr_offset = offsetof(struct backing_file, bf_freeptr);
bfilp_cachep = kmem_cache_create("bfilp", sizeof(struct backing_file),
&args, SLAB_HWCACHE_ALIGN | SLAB_PANIC |
SLAB_ACCOUNT | SLAB_TYPESAFE_BY_RCU);
percpu_counter_init(&nr_files, 0, GFP_KERNEL);
}
/*
* One file with associated inode and dcache is very roughly 1K. Per default
* do not use more than 10% of our memory for files.
*/
void __init files_maxfiles_init(void)
{
unsigned long n;
unsigned long nr_pages = totalram_pages();
mm: reference totalram_pages and managed_pages once per function Patch series "mm: convert totalram_pages, totalhigh_pages and managed pages to atomic", v5. This series converts totalram_pages, totalhigh_pages and zone->managed_pages to atomic variables. totalram_pages, zone->managed_pages and totalhigh_pages updates are protected by managed_page_count_lock, but readers never care about it. Convert these variables to atomic to avoid readers potentially seeing a store tear. Main motivation was that managed_page_count_lock handling was complicating things. It was discussed in length here, https://lore.kernel.org/patchwork/patch/995739/#1181785 It seemes better to remove the lock and convert variables to atomic. With the change, preventing poteintial store-to-read tearing comes as a bonus. This patch (of 4): This is in preparation to a later patch which converts totalram_pages and zone->managed_pages to atomic variables. Please note that re-reading the value might lead to a different value and as such it could lead to unexpected behavior. There are no known bugs as a result of the current code but it is better to prevent from them in principle. Link: http://lkml.kernel.org/r/1542090790-21750-2-git-send-email-arunks@codeaurora.org Signed-off-by: Arun KS <arunks@codeaurora.org> Reviewed-by: Konstantin Khlebnikov <khlebnikov@yandex-team.ru> Reviewed-by: David Hildenbrand <david@redhat.com> Acked-by: Michal Hocko <mhocko@suse.com> Acked-by: Vlastimil Babka <vbabka@suse.cz> Reviewed-by: Pavel Tatashin <pasha.tatashin@soleen.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2018-12-28 08:34:20 +00:00
unsigned long memreserve = (nr_pages - nr_free_pages()) * 3/2;
mm: reference totalram_pages and managed_pages once per function Patch series "mm: convert totalram_pages, totalhigh_pages and managed pages to atomic", v5. This series converts totalram_pages, totalhigh_pages and zone->managed_pages to atomic variables. totalram_pages, zone->managed_pages and totalhigh_pages updates are protected by managed_page_count_lock, but readers never care about it. Convert these variables to atomic to avoid readers potentially seeing a store tear. Main motivation was that managed_page_count_lock handling was complicating things. It was discussed in length here, https://lore.kernel.org/patchwork/patch/995739/#1181785 It seemes better to remove the lock and convert variables to atomic. With the change, preventing poteintial store-to-read tearing comes as a bonus. This patch (of 4): This is in preparation to a later patch which converts totalram_pages and zone->managed_pages to atomic variables. Please note that re-reading the value might lead to a different value and as such it could lead to unexpected behavior. There are no known bugs as a result of the current code but it is better to prevent from them in principle. Link: http://lkml.kernel.org/r/1542090790-21750-2-git-send-email-arunks@codeaurora.org Signed-off-by: Arun KS <arunks@codeaurora.org> Reviewed-by: Konstantin Khlebnikov <khlebnikov@yandex-team.ru> Reviewed-by: David Hildenbrand <david@redhat.com> Acked-by: Michal Hocko <mhocko@suse.com> Acked-by: Vlastimil Babka <vbabka@suse.cz> Reviewed-by: Pavel Tatashin <pasha.tatashin@soleen.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2018-12-28 08:34:20 +00:00
memreserve = min(memreserve, nr_pages - 1);
n = ((nr_pages - memreserve) * (PAGE_SIZE / 1024)) / 10;
fs: allow for more than 2^31 files Robin Holt tried to boot a 16TB system and found af_unix was overflowing a 32bit value : <quote> We were seeing a failure which prevented boot. The kernel was incapable of creating either a named pipe or unix domain socket. This comes down to a common kernel function called unix_create1() which does: atomic_inc(&unix_nr_socks); if (atomic_read(&unix_nr_socks) > 2 * get_max_files()) goto out; The function get_max_files() is a simple return of files_stat.max_files. files_stat.max_files is a signed integer and is computed in fs/file_table.c's files_init(). n = (mempages * (PAGE_SIZE / 1024)) / 10; files_stat.max_files = n; In our case, mempages (total_ram_pages) is approx 3,758,096,384 (0xe0000000). That leaves max_files at approximately 1,503,238,553. This causes 2 * get_max_files() to integer overflow. </quote> Fix is to let /proc/sys/fs/file-nr & /proc/sys/fs/file-max use long integers, and change af_unix to use an atomic_long_t instead of atomic_t. get_max_files() is changed to return an unsigned long. get_nr_files() is changed to return a long. unix_nr_socks is changed from atomic_t to atomic_long_t, while not strictly needed to address Robin problem. Before patch (on a 64bit kernel) : # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max -18446744071562067968 After patch: # echo 2147483648 >/proc/sys/fs/file-max # cat /proc/sys/fs/file-max 2147483648 # cat /proc/sys/fs/file-nr 704 0 2147483648 Reported-by: Robin Holt <holt@sgi.com> Signed-off-by: Eric Dumazet <eric.dumazet@gmail.com> Acked-by: David Miller <davem@davemloft.net> Reviewed-by: Robin Holt <holt@sgi.com> Tested-by: Robin Holt <holt@sgi.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Christoph Hellwig <hch@lst.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-10-26 21:22:44 +00:00
files_stat.max_files = max_t(unsigned long, n, NR_FILE);
}