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d9f68777b2
When xlog_sync() rounds off the tail the iclog that is being flushed, it manually subtracts that space from the grant heads. This space is actually reserved by the transaction ticket that covers the xlog_sync() call from xlog_write(), but we don't plumb the ticket down far enough for it to account for the space consumed in the current log ticket. The grant heads are hot, so we really should be accounting this to the ticket is we can, rather than adding thousands of extra grant head updates every CIL commit. Interestingly, this actually indicates a potential log space overrun can occur when we force the log. By the time that xfs_log_force() pushes out an active iclog and consumes the roundoff space, the reservation for that roundoff space has been returned to the grant heads and is no longer covered by a reservation. In theory the roundoff added to log force on an already full log could push the write head past the tail. In practice, the CIL commit that writes to the log and needs the iclog pushed will have reserved space for roundoff, so when it releases the ticket there will still be physical space for the roundoff to be committed to the log, even though it is no longer reserved. This roundoff won't be enough space to allow a transaction to be woken if the log is full, so overruns should not actually occur in practice. That said, it indicates that we should not release the CIL context log ticket until after we've released the commit iclog. It also means that xlog_sync() still needs the direct grant head manipulation if we don't provide it with a ticket. Log forces are rare when we are in fast paths running 1.5 million transactions/s that make the grant heads hot, so let's optimise the hot case and pass CIL log tickets down to the xlog_sync() code. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org>
714 lines
26 KiB
C
714 lines
26 KiB
C
// SPDX-License-Identifier: GPL-2.0
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/*
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* Copyright (c) 2000-2003,2005 Silicon Graphics, Inc.
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* All Rights Reserved.
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*/
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#ifndef __XFS_LOG_PRIV_H__
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#define __XFS_LOG_PRIV_H__
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struct xfs_buf;
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struct xlog;
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struct xlog_ticket;
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struct xfs_mount;
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/*
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* get client id from packed copy.
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*
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* this hack is here because the xlog_pack code copies four bytes
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* of xlog_op_header containing the fields oh_clientid, oh_flags
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* and oh_res2 into the packed copy.
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*
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* later on this four byte chunk is treated as an int and the
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* client id is pulled out.
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*
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* this has endian issues, of course.
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*/
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static inline uint xlog_get_client_id(__be32 i)
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{
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return be32_to_cpu(i) >> 24;
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}
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/*
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* In core log state
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*/
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enum xlog_iclog_state {
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XLOG_STATE_ACTIVE, /* Current IC log being written to */
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XLOG_STATE_WANT_SYNC, /* Want to sync this iclog; no more writes */
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XLOG_STATE_SYNCING, /* This IC log is syncing */
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XLOG_STATE_DONE_SYNC, /* Done syncing to disk */
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XLOG_STATE_CALLBACK, /* Callback functions now */
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XLOG_STATE_DIRTY, /* Dirty IC log, not ready for ACTIVE status */
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};
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#define XLOG_STATE_STRINGS \
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{ XLOG_STATE_ACTIVE, "XLOG_STATE_ACTIVE" }, \
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{ XLOG_STATE_WANT_SYNC, "XLOG_STATE_WANT_SYNC" }, \
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{ XLOG_STATE_SYNCING, "XLOG_STATE_SYNCING" }, \
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{ XLOG_STATE_DONE_SYNC, "XLOG_STATE_DONE_SYNC" }, \
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{ XLOG_STATE_CALLBACK, "XLOG_STATE_CALLBACK" }, \
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{ XLOG_STATE_DIRTY, "XLOG_STATE_DIRTY" }
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/*
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* In core log flags
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*/
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#define XLOG_ICL_NEED_FLUSH (1u << 0) /* iclog needs REQ_PREFLUSH */
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#define XLOG_ICL_NEED_FUA (1u << 1) /* iclog needs REQ_FUA */
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#define XLOG_ICL_STRINGS \
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{ XLOG_ICL_NEED_FLUSH, "XLOG_ICL_NEED_FLUSH" }, \
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{ XLOG_ICL_NEED_FUA, "XLOG_ICL_NEED_FUA" }
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/*
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* Log ticket flags
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*/
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#define XLOG_TIC_PERM_RESERV (1u << 0) /* permanent reservation */
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#define XLOG_TIC_FLAGS \
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{ XLOG_TIC_PERM_RESERV, "XLOG_TIC_PERM_RESERV" }
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/*
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* Below are states for covering allocation transactions.
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* By covering, we mean changing the h_tail_lsn in the last on-disk
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* log write such that no allocation transactions will be re-done during
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* recovery after a system crash. Recovery starts at the last on-disk
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* log write.
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*
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* These states are used to insert dummy log entries to cover
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* space allocation transactions which can undo non-transactional changes
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* after a crash. Writes to a file with space
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* already allocated do not result in any transactions. Allocations
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* might include space beyond the EOF. So if we just push the EOF a
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* little, the last transaction for the file could contain the wrong
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* size. If there is no file system activity, after an allocation
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* transaction, and the system crashes, the allocation transaction
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* will get replayed and the file will be truncated. This could
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* be hours/days/... after the allocation occurred.
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*
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* The fix for this is to do two dummy transactions when the
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* system is idle. We need two dummy transaction because the h_tail_lsn
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* in the log record header needs to point beyond the last possible
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* non-dummy transaction. The first dummy changes the h_tail_lsn to
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* the first transaction before the dummy. The second dummy causes
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* h_tail_lsn to point to the first dummy. Recovery starts at h_tail_lsn.
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*
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* These dummy transactions get committed when everything
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* is idle (after there has been some activity).
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*
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* There are 5 states used to control this.
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*
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* IDLE -- no logging has been done on the file system or
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* we are done covering previous transactions.
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* NEED -- logging has occurred and we need a dummy transaction
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* when the log becomes idle.
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* DONE -- we were in the NEED state and have committed a dummy
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* transaction.
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* NEED2 -- we detected that a dummy transaction has gone to the
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* on disk log with no other transactions.
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* DONE2 -- we committed a dummy transaction when in the NEED2 state.
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*
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* There are two places where we switch states:
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*
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* 1.) In xfs_sync, when we detect an idle log and are in NEED or NEED2.
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* We commit the dummy transaction and switch to DONE or DONE2,
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* respectively. In all other states, we don't do anything.
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*
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* 2.) When we finish writing the on-disk log (xlog_state_clean_log).
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*
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* No matter what state we are in, if this isn't the dummy
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* transaction going out, the next state is NEED.
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* So, if we aren't in the DONE or DONE2 states, the next state
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* is NEED. We can't be finishing a write of the dummy record
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* unless it was committed and the state switched to DONE or DONE2.
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*
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* If we are in the DONE state and this was a write of the
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* dummy transaction, we move to NEED2.
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*
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* If we are in the DONE2 state and this was a write of the
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* dummy transaction, we move to IDLE.
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*
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*
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* Writing only one dummy transaction can get appended to
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* one file space allocation. When this happens, the log recovery
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* code replays the space allocation and a file could be truncated.
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* This is why we have the NEED2 and DONE2 states before going idle.
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*/
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#define XLOG_STATE_COVER_IDLE 0
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#define XLOG_STATE_COVER_NEED 1
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#define XLOG_STATE_COVER_DONE 2
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#define XLOG_STATE_COVER_NEED2 3
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#define XLOG_STATE_COVER_DONE2 4
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#define XLOG_COVER_OPS 5
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typedef struct xlog_ticket {
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struct list_head t_queue; /* reserve/write queue */
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struct task_struct *t_task; /* task that owns this ticket */
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xlog_tid_t t_tid; /* transaction identifier */
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atomic_t t_ref; /* ticket reference count */
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int t_curr_res; /* current reservation */
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int t_unit_res; /* unit reservation */
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char t_ocnt; /* original unit count */
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char t_cnt; /* current unit count */
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uint8_t t_flags; /* properties of reservation */
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int t_iclog_hdrs; /* iclog hdrs in t_curr_res */
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} xlog_ticket_t;
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/*
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* - A log record header is 512 bytes. There is plenty of room to grow the
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* xlog_rec_header_t into the reserved space.
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* - ic_data follows, so a write to disk can start at the beginning of
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* the iclog.
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* - ic_forcewait is used to implement synchronous forcing of the iclog to disk.
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* - ic_next is the pointer to the next iclog in the ring.
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* - ic_log is a pointer back to the global log structure.
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* - ic_size is the full size of the log buffer, minus the cycle headers.
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* - ic_offset is the current number of bytes written to in this iclog.
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* - ic_refcnt is bumped when someone is writing to the log.
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* - ic_state is the state of the iclog.
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*
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* Because of cacheline contention on large machines, we need to separate
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* various resources onto different cachelines. To start with, make the
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* structure cacheline aligned. The following fields can be contended on
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* by independent processes:
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*
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* - ic_callbacks
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* - ic_refcnt
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* - fields protected by the global l_icloglock
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*
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* so we need to ensure that these fields are located in separate cachelines.
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* We'll put all the read-only and l_icloglock fields in the first cacheline,
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* and move everything else out to subsequent cachelines.
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*/
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typedef struct xlog_in_core {
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wait_queue_head_t ic_force_wait;
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wait_queue_head_t ic_write_wait;
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struct xlog_in_core *ic_next;
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struct xlog_in_core *ic_prev;
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struct xlog *ic_log;
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u32 ic_size;
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u32 ic_offset;
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enum xlog_iclog_state ic_state;
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unsigned int ic_flags;
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void *ic_datap; /* pointer to iclog data */
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struct list_head ic_callbacks;
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/* reference counts need their own cacheline */
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atomic_t ic_refcnt ____cacheline_aligned_in_smp;
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xlog_in_core_2_t *ic_data;
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#define ic_header ic_data->hic_header
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#ifdef DEBUG
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bool ic_fail_crc : 1;
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#endif
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struct semaphore ic_sema;
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struct work_struct ic_end_io_work;
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struct bio ic_bio;
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struct bio_vec ic_bvec[];
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} xlog_in_core_t;
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/*
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* The CIL context is used to aggregate per-transaction details as well be
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* passed to the iclog for checkpoint post-commit processing. After being
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* passed to the iclog, another context needs to be allocated for tracking the
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* next set of transactions to be aggregated into a checkpoint.
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*/
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struct xfs_cil;
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struct xfs_cil_ctx {
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struct xfs_cil *cil;
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xfs_csn_t sequence; /* chkpt sequence # */
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xfs_lsn_t start_lsn; /* first LSN of chkpt commit */
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xfs_lsn_t commit_lsn; /* chkpt commit record lsn */
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struct xlog_in_core *commit_iclog;
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struct xlog_ticket *ticket; /* chkpt ticket */
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atomic_t space_used; /* aggregate size of regions */
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struct list_head busy_extents; /* busy extents in chkpt */
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struct list_head log_items; /* log items in chkpt */
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struct list_head lv_chain; /* logvecs being pushed */
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struct list_head iclog_entry;
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struct list_head committing; /* ctx committing list */
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struct work_struct discard_endio_work;
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struct work_struct push_work;
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atomic_t order_id;
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};
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/*
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* Per-cpu CIL tracking items
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*/
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struct xlog_cil_pcp {
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int32_t space_used;
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uint32_t space_reserved;
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struct list_head busy_extents;
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struct list_head log_items;
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};
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/*
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* Committed Item List structure
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*
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* This structure is used to track log items that have been committed but not
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* yet written into the log. It is used only when the delayed logging mount
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* option is enabled.
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*
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* This structure tracks the list of committing checkpoint contexts so
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* we can avoid the problem of having to hold out new transactions during a
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* flush until we have a the commit record LSN of the checkpoint. We can
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* traverse the list of committing contexts in xlog_cil_push_lsn() to find a
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* sequence match and extract the commit LSN directly from there. If the
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* checkpoint is still in the process of committing, we can block waiting for
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* the commit LSN to be determined as well. This should make synchronous
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* operations almost as efficient as the old logging methods.
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*/
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struct xfs_cil {
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struct xlog *xc_log;
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unsigned long xc_flags;
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atomic_t xc_iclog_hdrs;
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struct workqueue_struct *xc_push_wq;
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struct rw_semaphore xc_ctx_lock ____cacheline_aligned_in_smp;
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struct xfs_cil_ctx *xc_ctx;
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spinlock_t xc_push_lock ____cacheline_aligned_in_smp;
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xfs_csn_t xc_push_seq;
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bool xc_push_commit_stable;
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struct list_head xc_committing;
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wait_queue_head_t xc_commit_wait;
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wait_queue_head_t xc_start_wait;
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xfs_csn_t xc_current_sequence;
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wait_queue_head_t xc_push_wait; /* background push throttle */
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void __percpu *xc_pcp; /* percpu CIL structures */
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#ifdef CONFIG_HOTPLUG_CPU
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struct list_head xc_pcp_list;
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#endif
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} ____cacheline_aligned_in_smp;
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/* xc_flags bit values */
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#define XLOG_CIL_EMPTY 1
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#define XLOG_CIL_PCP_SPACE 2
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/*
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* The amount of log space we allow the CIL to aggregate is difficult to size.
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* Whatever we choose, we have to make sure we can get a reservation for the
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* log space effectively, that it is large enough to capture sufficient
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* relogging to reduce log buffer IO significantly, but it is not too large for
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* the log or induces too much latency when writing out through the iclogs. We
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* track both space consumed and the number of vectors in the checkpoint
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* context, so we need to decide which to use for limiting.
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*
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* Every log buffer we write out during a push needs a header reserved, which
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* is at least one sector and more for v2 logs. Hence we need a reservation of
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* at least 512 bytes per 32k of log space just for the LR headers. That means
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* 16KB of reservation per megabyte of delayed logging space we will consume,
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* plus various headers. The number of headers will vary based on the num of
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* io vectors, so limiting on a specific number of vectors is going to result
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* in transactions of varying size. IOWs, it is more consistent to track and
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* limit space consumed in the log rather than by the number of objects being
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* logged in order to prevent checkpoint ticket overruns.
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*
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* Further, use of static reservations through the log grant mechanism is
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* problematic. It introduces a lot of complexity (e.g. reserve grant vs write
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* grant) and a significant deadlock potential because regranting write space
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* can block on log pushes. Hence if we have to regrant log space during a log
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* push, we can deadlock.
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*
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* However, we can avoid this by use of a dynamic "reservation stealing"
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* technique during transaction commit whereby unused reservation space in the
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* transaction ticket is transferred to the CIL ctx commit ticket to cover the
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* space needed by the checkpoint transaction. This means that we never need to
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* specifically reserve space for the CIL checkpoint transaction, nor do we
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* need to regrant space once the checkpoint completes. This also means the
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* checkpoint transaction ticket is specific to the checkpoint context, rather
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* than the CIL itself.
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*
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* With dynamic reservations, we can effectively make up arbitrary limits for
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* the checkpoint size so long as they don't violate any other size rules.
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* Recovery imposes a rule that no transaction exceed half the log, so we are
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* limited by that. Furthermore, the log transaction reservation subsystem
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* tries to keep 25% of the log free, so we need to keep below that limit or we
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* risk running out of free log space to start any new transactions.
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*
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* In order to keep background CIL push efficient, we only need to ensure the
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* CIL is large enough to maintain sufficient in-memory relogging to avoid
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* repeated physical writes of frequently modified metadata. If we allow the CIL
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* to grow to a substantial fraction of the log, then we may be pinning hundreds
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* of megabytes of metadata in memory until the CIL flushes. This can cause
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* issues when we are running low on memory - pinned memory cannot be reclaimed,
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* and the CIL consumes a lot of memory. Hence we need to set an upper physical
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* size limit for the CIL that limits the maximum amount of memory pinned by the
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* CIL but does not limit performance by reducing relogging efficiency
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* significantly.
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*
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* As such, the CIL push threshold ends up being the smaller of two thresholds:
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* - a threshold large enough that it allows CIL to be pushed and progress to be
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* made without excessive blocking of incoming transaction commits. This is
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* defined to be 12.5% of the log space - half the 25% push threshold of the
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* AIL.
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* - small enough that it doesn't pin excessive amounts of memory but maintains
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* close to peak relogging efficiency. This is defined to be 16x the iclog
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* buffer window (32MB) as measurements have shown this to be roughly the
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* point of diminishing performance increases under highly concurrent
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* modification workloads.
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*
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* To prevent the CIL from overflowing upper commit size bounds, we introduce a
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* new threshold at which we block committing transactions until the background
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* CIL commit commences and switches to a new context. While this is not a hard
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* limit, it forces the process committing a transaction to the CIL to block and
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* yeild the CPU, giving the CIL push work a chance to be scheduled and start
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* work. This prevents a process running lots of transactions from overfilling
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* the CIL because it is not yielding the CPU. We set the blocking limit at
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* twice the background push space threshold so we keep in line with the AIL
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* push thresholds.
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*
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* Note: this is not a -hard- limit as blocking is applied after the transaction
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* is inserted into the CIL and the push has been triggered. It is largely a
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* throttling mechanism that allows the CIL push to be scheduled and run. A hard
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* limit will be difficult to implement without introducing global serialisation
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* in the CIL commit fast path, and it's not at all clear that we actually need
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* such hard limits given the ~7 years we've run without a hard limit before
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* finding the first situation where a checkpoint size overflow actually
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* occurred. Hence the simple throttle, and an ASSERT check to tell us that
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* we've overrun the max size.
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*/
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#define XLOG_CIL_SPACE_LIMIT(log) \
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min_t(int, (log)->l_logsize >> 3, BBTOB(XLOG_TOTAL_REC_SHIFT(log)) << 4)
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#define XLOG_CIL_BLOCKING_SPACE_LIMIT(log) \
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(XLOG_CIL_SPACE_LIMIT(log) * 2)
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/*
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* ticket grant locks, queues and accounting have their own cachlines
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* as these are quite hot and can be operated on concurrently.
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*/
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struct xlog_grant_head {
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spinlock_t lock ____cacheline_aligned_in_smp;
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struct list_head waiters;
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atomic64_t grant;
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};
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/*
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* The reservation head lsn is not made up of a cycle number and block number.
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* Instead, it uses a cycle number and byte number. Logs don't expect to
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* overflow 31 bits worth of byte offset, so using a byte number will mean
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* that round off problems won't occur when releasing partial reservations.
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*/
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struct xlog {
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/* The following fields don't need locking */
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struct xfs_mount *l_mp; /* mount point */
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struct xfs_ail *l_ailp; /* AIL log is working with */
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struct xfs_cil *l_cilp; /* CIL log is working with */
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struct xfs_buftarg *l_targ; /* buftarg of log */
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struct workqueue_struct *l_ioend_workqueue; /* for I/O completions */
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struct delayed_work l_work; /* background flush work */
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long l_opstate; /* operational state */
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uint l_quotaoffs_flag; /* XFS_DQ_*, for QUOTAOFFs */
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struct list_head *l_buf_cancel_table;
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int l_iclog_hsize; /* size of iclog header */
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int l_iclog_heads; /* # of iclog header sectors */
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uint l_sectBBsize; /* sector size in BBs (2^n) */
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int l_iclog_size; /* size of log in bytes */
|
|
int l_iclog_bufs; /* number of iclog buffers */
|
|
xfs_daddr_t l_logBBstart; /* start block of log */
|
|
int l_logsize; /* size of log in bytes */
|
|
int l_logBBsize; /* size of log in BB chunks */
|
|
|
|
/* The following block of fields are changed while holding icloglock */
|
|
wait_queue_head_t l_flush_wait ____cacheline_aligned_in_smp;
|
|
/* waiting for iclog flush */
|
|
int l_covered_state;/* state of "covering disk
|
|
* log entries" */
|
|
xlog_in_core_t *l_iclog; /* head log queue */
|
|
spinlock_t l_icloglock; /* grab to change iclog state */
|
|
int l_curr_cycle; /* Cycle number of log writes */
|
|
int l_prev_cycle; /* Cycle number before last
|
|
* block increment */
|
|
int l_curr_block; /* current logical log block */
|
|
int l_prev_block; /* previous logical log block */
|
|
|
|
/*
|
|
* l_last_sync_lsn and l_tail_lsn are atomics so they can be set and
|
|
* read without needing to hold specific locks. To avoid operations
|
|
* contending with other hot objects, place each of them on a separate
|
|
* cacheline.
|
|
*/
|
|
/* lsn of last LR on disk */
|
|
atomic64_t l_last_sync_lsn ____cacheline_aligned_in_smp;
|
|
/* lsn of 1st LR with unflushed * buffers */
|
|
atomic64_t l_tail_lsn ____cacheline_aligned_in_smp;
|
|
|
|
struct xlog_grant_head l_reserve_head;
|
|
struct xlog_grant_head l_write_head;
|
|
|
|
struct xfs_kobj l_kobj;
|
|
|
|
/* log recovery lsn tracking (for buffer submission */
|
|
xfs_lsn_t l_recovery_lsn;
|
|
|
|
uint32_t l_iclog_roundoff;/* padding roundoff */
|
|
|
|
/* Users of log incompat features should take a read lock. */
|
|
struct rw_semaphore l_incompat_users;
|
|
};
|
|
|
|
/*
|
|
* Bits for operational state
|
|
*/
|
|
#define XLOG_ACTIVE_RECOVERY 0 /* in the middle of recovery */
|
|
#define XLOG_RECOVERY_NEEDED 1 /* log was recovered */
|
|
#define XLOG_IO_ERROR 2 /* log hit an I/O error, and being
|
|
shutdown */
|
|
#define XLOG_TAIL_WARN 3 /* log tail verify warning issued */
|
|
|
|
static inline bool
|
|
xlog_recovery_needed(struct xlog *log)
|
|
{
|
|
return test_bit(XLOG_RECOVERY_NEEDED, &log->l_opstate);
|
|
}
|
|
|
|
static inline bool
|
|
xlog_in_recovery(struct xlog *log)
|
|
{
|
|
return test_bit(XLOG_ACTIVE_RECOVERY, &log->l_opstate);
|
|
}
|
|
|
|
static inline bool
|
|
xlog_is_shutdown(struct xlog *log)
|
|
{
|
|
return test_bit(XLOG_IO_ERROR, &log->l_opstate);
|
|
}
|
|
|
|
/*
|
|
* Wait until the xlog_force_shutdown() has marked the log as shut down
|
|
* so xlog_is_shutdown() will always return true.
|
|
*/
|
|
static inline void
|
|
xlog_shutdown_wait(
|
|
struct xlog *log)
|
|
{
|
|
wait_var_event(&log->l_opstate, xlog_is_shutdown(log));
|
|
}
|
|
|
|
/* common routines */
|
|
extern int
|
|
xlog_recover(
|
|
struct xlog *log);
|
|
extern int
|
|
xlog_recover_finish(
|
|
struct xlog *log);
|
|
extern void
|
|
xlog_recover_cancel(struct xlog *);
|
|
|
|
extern __le32 xlog_cksum(struct xlog *log, struct xlog_rec_header *rhead,
|
|
char *dp, int size);
|
|
|
|
extern struct kmem_cache *xfs_log_ticket_cache;
|
|
struct xlog_ticket *xlog_ticket_alloc(struct xlog *log, int unit_bytes,
|
|
int count, bool permanent);
|
|
|
|
void xlog_print_tic_res(struct xfs_mount *mp, struct xlog_ticket *ticket);
|
|
void xlog_print_trans(struct xfs_trans *);
|
|
int xlog_write(struct xlog *log, struct xfs_cil_ctx *ctx,
|
|
struct list_head *lv_chain, struct xlog_ticket *tic,
|
|
uint32_t len);
|
|
void xfs_log_ticket_ungrant(struct xlog *log, struct xlog_ticket *ticket);
|
|
void xfs_log_ticket_regrant(struct xlog *log, struct xlog_ticket *ticket);
|
|
|
|
void xlog_state_switch_iclogs(struct xlog *log, struct xlog_in_core *iclog,
|
|
int eventual_size);
|
|
int xlog_state_release_iclog(struct xlog *log, struct xlog_in_core *iclog,
|
|
struct xlog_ticket *ticket);
|
|
|
|
/*
|
|
* When we crack an atomic LSN, we sample it first so that the value will not
|
|
* change while we are cracking it into the component values. This means we
|
|
* will always get consistent component values to work from. This should always
|
|
* be used to sample and crack LSNs that are stored and updated in atomic
|
|
* variables.
|
|
*/
|
|
static inline void
|
|
xlog_crack_atomic_lsn(atomic64_t *lsn, uint *cycle, uint *block)
|
|
{
|
|
xfs_lsn_t val = atomic64_read(lsn);
|
|
|
|
*cycle = CYCLE_LSN(val);
|
|
*block = BLOCK_LSN(val);
|
|
}
|
|
|
|
/*
|
|
* Calculate and assign a value to an atomic LSN variable from component pieces.
|
|
*/
|
|
static inline void
|
|
xlog_assign_atomic_lsn(atomic64_t *lsn, uint cycle, uint block)
|
|
{
|
|
atomic64_set(lsn, xlog_assign_lsn(cycle, block));
|
|
}
|
|
|
|
/*
|
|
* When we crack the grant head, we sample it first so that the value will not
|
|
* change while we are cracking it into the component values. This means we
|
|
* will always get consistent component values to work from.
|
|
*/
|
|
static inline void
|
|
xlog_crack_grant_head_val(int64_t val, int *cycle, int *space)
|
|
{
|
|
*cycle = val >> 32;
|
|
*space = val & 0xffffffff;
|
|
}
|
|
|
|
static inline void
|
|
xlog_crack_grant_head(atomic64_t *head, int *cycle, int *space)
|
|
{
|
|
xlog_crack_grant_head_val(atomic64_read(head), cycle, space);
|
|
}
|
|
|
|
static inline int64_t
|
|
xlog_assign_grant_head_val(int cycle, int space)
|
|
{
|
|
return ((int64_t)cycle << 32) | space;
|
|
}
|
|
|
|
static inline void
|
|
xlog_assign_grant_head(atomic64_t *head, int cycle, int space)
|
|
{
|
|
atomic64_set(head, xlog_assign_grant_head_val(cycle, space));
|
|
}
|
|
|
|
/*
|
|
* Committed Item List interfaces
|
|
*/
|
|
int xlog_cil_init(struct xlog *log);
|
|
void xlog_cil_init_post_recovery(struct xlog *log);
|
|
void xlog_cil_destroy(struct xlog *log);
|
|
bool xlog_cil_empty(struct xlog *log);
|
|
void xlog_cil_commit(struct xlog *log, struct xfs_trans *tp,
|
|
xfs_csn_t *commit_seq, bool regrant);
|
|
void xlog_cil_set_ctx_write_state(struct xfs_cil_ctx *ctx,
|
|
struct xlog_in_core *iclog);
|
|
|
|
|
|
/*
|
|
* CIL force routines
|
|
*/
|
|
void xlog_cil_flush(struct xlog *log);
|
|
xfs_lsn_t xlog_cil_force_seq(struct xlog *log, xfs_csn_t sequence);
|
|
|
|
static inline void
|
|
xlog_cil_force(struct xlog *log)
|
|
{
|
|
xlog_cil_force_seq(log, log->l_cilp->xc_current_sequence);
|
|
}
|
|
|
|
/*
|
|
* Wrapper function for waiting on a wait queue serialised against wakeups
|
|
* by a spinlock. This matches the semantics of all the wait queues used in the
|
|
* log code.
|
|
*/
|
|
static inline void
|
|
xlog_wait(
|
|
struct wait_queue_head *wq,
|
|
struct spinlock *lock)
|
|
__releases(lock)
|
|
{
|
|
DECLARE_WAITQUEUE(wait, current);
|
|
|
|
add_wait_queue_exclusive(wq, &wait);
|
|
__set_current_state(TASK_UNINTERRUPTIBLE);
|
|
spin_unlock(lock);
|
|
schedule();
|
|
remove_wait_queue(wq, &wait);
|
|
}
|
|
|
|
int xlog_wait_on_iclog(struct xlog_in_core *iclog);
|
|
|
|
/*
|
|
* The LSN is valid so long as it is behind the current LSN. If it isn't, this
|
|
* means that the next log record that includes this metadata could have a
|
|
* smaller LSN. In turn, this means that the modification in the log would not
|
|
* replay.
|
|
*/
|
|
static inline bool
|
|
xlog_valid_lsn(
|
|
struct xlog *log,
|
|
xfs_lsn_t lsn)
|
|
{
|
|
int cur_cycle;
|
|
int cur_block;
|
|
bool valid = true;
|
|
|
|
/*
|
|
* First, sample the current lsn without locking to avoid added
|
|
* contention from metadata I/O. The current cycle and block are updated
|
|
* (in xlog_state_switch_iclogs()) and read here in a particular order
|
|
* to avoid false negatives (e.g., thinking the metadata LSN is valid
|
|
* when it is not).
|
|
*
|
|
* The current block is always rewound before the cycle is bumped in
|
|
* xlog_state_switch_iclogs() to ensure the current LSN is never seen in
|
|
* a transiently forward state. Instead, we can see the LSN in a
|
|
* transiently behind state if we happen to race with a cycle wrap.
|
|
*/
|
|
cur_cycle = READ_ONCE(log->l_curr_cycle);
|
|
smp_rmb();
|
|
cur_block = READ_ONCE(log->l_curr_block);
|
|
|
|
if ((CYCLE_LSN(lsn) > cur_cycle) ||
|
|
(CYCLE_LSN(lsn) == cur_cycle && BLOCK_LSN(lsn) > cur_block)) {
|
|
/*
|
|
* If the metadata LSN appears invalid, it's possible the check
|
|
* above raced with a wrap to the next log cycle. Grab the lock
|
|
* to check for sure.
|
|
*/
|
|
spin_lock(&log->l_icloglock);
|
|
cur_cycle = log->l_curr_cycle;
|
|
cur_block = log->l_curr_block;
|
|
spin_unlock(&log->l_icloglock);
|
|
|
|
if ((CYCLE_LSN(lsn) > cur_cycle) ||
|
|
(CYCLE_LSN(lsn) == cur_cycle && BLOCK_LSN(lsn) > cur_block))
|
|
valid = false;
|
|
}
|
|
|
|
return valid;
|
|
}
|
|
|
|
/*
|
|
* Log vector and shadow buffers can be large, so we need to use kvmalloc() here
|
|
* to ensure success. Unfortunately, kvmalloc() only allows GFP_KERNEL contexts
|
|
* to fall back to vmalloc, so we can't actually do anything useful with gfp
|
|
* flags to control the kmalloc() behaviour within kvmalloc(). Hence kmalloc()
|
|
* will do direct reclaim and compaction in the slow path, both of which are
|
|
* horrendously expensive. We just want kmalloc to fail fast and fall back to
|
|
* vmalloc if it can't get somethign straight away from the free lists or
|
|
* buddy allocator. Hence we have to open code kvmalloc outselves here.
|
|
*
|
|
* This assumes that the caller uses memalloc_nofs_save task context here, so
|
|
* despite the use of GFP_KERNEL here, we are going to be doing GFP_NOFS
|
|
* allocations. This is actually the only way to make vmalloc() do GFP_NOFS
|
|
* allocations, so lets just all pretend this is a GFP_KERNEL context
|
|
* operation....
|
|
*/
|
|
static inline void *
|
|
xlog_kvmalloc(
|
|
size_t buf_size)
|
|
{
|
|
gfp_t flags = GFP_KERNEL;
|
|
void *p;
|
|
|
|
flags &= ~__GFP_DIRECT_RECLAIM;
|
|
flags |= __GFP_NOWARN | __GFP_NORETRY;
|
|
do {
|
|
p = kmalloc(buf_size, flags);
|
|
if (!p)
|
|
p = vmalloc(buf_size);
|
|
} while (!p);
|
|
|
|
return p;
|
|
}
|
|
|
|
/*
|
|
* CIL CPU dead notifier
|
|
*/
|
|
void xlog_cil_pcp_dead(struct xlog *log, unsigned int cpu);
|
|
|
|
#endif /* __XFS_LOG_PRIV_H__ */
|