OpenCloudOS-Kernel/fs/xfs/xfs_file.c

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// SPDX-License-Identifier: GPL-2.0
/*
* Copyright (c) 2000-2005 Silicon Graphics, Inc.
* All Rights Reserved.
*/
#include "xfs.h"
#include "xfs_fs.h"
#include "xfs_shared.h"
#include "xfs_format.h"
#include "xfs_log_format.h"
#include "xfs_trans_resv.h"
#include "xfs_mount.h"
#include "xfs_inode.h"
#include "xfs_trans.h"
#include "xfs_inode_item.h"
#include "xfs_bmap.h"
#include "xfs_bmap_util.h"
#include "xfs_dir2.h"
#include "xfs_dir2_priv.h"
#include "xfs_ioctl.h"
#include "xfs_trace.h"
#include "xfs_log.h"
xfs: run an eofblocks scan on ENOSPC/EDQUOT From: Brian Foster <bfoster@redhat.com> Speculative preallocation and and the associated throttling metrics assume we're working with large files on large filesystems. Users have reported inefficiencies in these mechanisms when we happen to be dealing with large files on smaller filesystems. This can occur because while prealloc throttling is aggressive under low free space conditions, it is not active until we reach 5% free space or less. For example, a 40GB filesystem has enough space for several files large enough to have multi-GB preallocations at any given time. If those files are slow growing, they might reserve preallocation for long periods of time as well as avoid the background scanner due to frequent modification. If a new file is written under these conditions, said file has no access to this already reserved space and premature ENOSPC is imminent. To handle this scenario, modify the buffered write ENOSPC handling and retry sequence to invoke an eofblocks scan. In the smaller filesystem scenario, the eofblocks scan resets the usage of preallocation such that when the 5% free space threshold is met, throttling effectively takes over to provide fair and efficient preallocation until legitimate ENOSPC. The eofblocks scan is selective based on the nature of the failure. For example, an EDQUOT failure in a particular quota will use a filtered scan for that quota. Because we don't know which quota might have caused an allocation failure at any given time, we include each applicable quota determined to be under low free space conditions in the scan. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2014-07-24 17:49:28 +08:00
#include "xfs_icache.h"
#include "xfs_pnfs.h"
#include "xfs_iomap.h"
#include "xfs_reflink.h"
#include <linux/dax.h>
#include <linux/falloc.h>
#include <linux/backing-dev.h>
#include <linux/mman.h>
#include <linux/fadvise.h>
#include <linux/mount.h>
static const struct vm_operations_struct xfs_file_vm_ops;
/*
* Decide if the given file range is aligned to the size of the fundamental
* allocation unit for the file.
*/
static bool
xfs_is_falloc_aligned(
struct xfs_inode *ip,
loff_t pos,
long long int len)
{
struct xfs_mount *mp = ip->i_mount;
uint64_t mask;
if (XFS_IS_REALTIME_INODE(ip)) {
if (!is_power_of_2(mp->m_sb.sb_rextsize)) {
u64 rextbytes;
u32 mod;
rextbytes = XFS_FSB_TO_B(mp, mp->m_sb.sb_rextsize);
div_u64_rem(pos, rextbytes, &mod);
if (mod)
return false;
div_u64_rem(len, rextbytes, &mod);
return mod == 0;
}
mask = XFS_FSB_TO_B(mp, mp->m_sb.sb_rextsize) - 1;
} else {
mask = mp->m_sb.sb_blocksize - 1;
}
return !((pos | len) & mask);
}
/*
* Fsync operations on directories are much simpler than on regular files,
* as there is no file data to flush, and thus also no need for explicit
* cache flush operations, and there are no non-transaction metadata updates
* on directories either.
*/
STATIC int
xfs_dir_fsync(
struct file *file,
loff_t start,
loff_t end,
int datasync)
{
struct xfs_inode *ip = XFS_I(file->f_mapping->host);
trace_xfs_dir_fsync(ip);
return xfs_log_force_inode(ip);
}
xfs: xfs_log_force_lsn isn't passed a LSN In doing an investigation into AIL push stalls, I was looking at the log force code to see if an async CIL push could be done instead. This lead me to xfs_log_force_lsn() and looking at how it works. xfs_log_force_lsn() is only called from inode synchronisation contexts such as fsync(), and it takes the ip->i_itemp->ili_last_lsn value as the LSN to sync the log to. This gets passed to xlog_cil_force_lsn() via xfs_log_force_lsn() to flush the CIL to the journal, and then used by xfs_log_force_lsn() to flush the iclogs to the journal. The problem is that ip->i_itemp->ili_last_lsn does not store a log sequence number. What it stores is passed to it from the ->iop_committing method, which is called by xfs_log_commit_cil(). The value this passes to the iop_committing method is the CIL context sequence number that the item was committed to. As it turns out, xlog_cil_force_lsn() converts the sequence to an actual commit LSN for the related context and returns that to xfs_log_force_lsn(). xfs_log_force_lsn() overwrites it's "lsn" variable that contained a sequence with an actual LSN and then uses that to sync the iclogs. This caused me some confusion for a while, even though I originally wrote all this code a decade ago. ->iop_committing is only used by a couple of log item types, and only inode items use the sequence number it is passed. Let's clean up the API, CIL structures and inode log item to call it a sequence number, and make it clear that the high level code is using CIL sequence numbers and not on-disk LSNs for integrity synchronisation purposes. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Allison Henderson <allison.henderson@oracle.com> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-06-18 23:21:52 +08:00
static xfs_csn_t
xfs_fsync_seq(
struct xfs_inode *ip,
bool datasync)
{
if (!xfs_ipincount(ip))
return 0;
if (datasync && !(ip->i_itemp->ili_fsync_fields & ~XFS_ILOG_TIMESTAMP))
return 0;
xfs: xfs_log_force_lsn isn't passed a LSN In doing an investigation into AIL push stalls, I was looking at the log force code to see if an async CIL push could be done instead. This lead me to xfs_log_force_lsn() and looking at how it works. xfs_log_force_lsn() is only called from inode synchronisation contexts such as fsync(), and it takes the ip->i_itemp->ili_last_lsn value as the LSN to sync the log to. This gets passed to xlog_cil_force_lsn() via xfs_log_force_lsn() to flush the CIL to the journal, and then used by xfs_log_force_lsn() to flush the iclogs to the journal. The problem is that ip->i_itemp->ili_last_lsn does not store a log sequence number. What it stores is passed to it from the ->iop_committing method, which is called by xfs_log_commit_cil(). The value this passes to the iop_committing method is the CIL context sequence number that the item was committed to. As it turns out, xlog_cil_force_lsn() converts the sequence to an actual commit LSN for the related context and returns that to xfs_log_force_lsn(). xfs_log_force_lsn() overwrites it's "lsn" variable that contained a sequence with an actual LSN and then uses that to sync the iclogs. This caused me some confusion for a while, even though I originally wrote all this code a decade ago. ->iop_committing is only used by a couple of log item types, and only inode items use the sequence number it is passed. Let's clean up the API, CIL structures and inode log item to call it a sequence number, and make it clear that the high level code is using CIL sequence numbers and not on-disk LSNs for integrity synchronisation purposes. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Allison Henderson <allison.henderson@oracle.com> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-06-18 23:21:52 +08:00
return ip->i_itemp->ili_commit_seq;
}
/*
* All metadata updates are logged, which means that we just have to flush the
* log up to the latest LSN that touched the inode.
*
* If we have concurrent fsync/fdatasync() calls, we need them to all block on
* the log force before we clear the ili_fsync_fields field. This ensures that
* we don't get a racing sync operation that does not wait for the metadata to
* hit the journal before returning. If we race with clearing ili_fsync_fields,
* then all that will happen is the log force will do nothing as the lsn will
* already be on disk. We can't race with setting ili_fsync_fields because that
* is done under XFS_ILOCK_EXCL, and that can't happen because we hold the lock
* shared until after the ili_fsync_fields is cleared.
*/
static int
xfs_fsync_flush_log(
struct xfs_inode *ip,
bool datasync,
int *log_flushed)
{
int error = 0;
xfs: xfs_log_force_lsn isn't passed a LSN In doing an investigation into AIL push stalls, I was looking at the log force code to see if an async CIL push could be done instead. This lead me to xfs_log_force_lsn() and looking at how it works. xfs_log_force_lsn() is only called from inode synchronisation contexts such as fsync(), and it takes the ip->i_itemp->ili_last_lsn value as the LSN to sync the log to. This gets passed to xlog_cil_force_lsn() via xfs_log_force_lsn() to flush the CIL to the journal, and then used by xfs_log_force_lsn() to flush the iclogs to the journal. The problem is that ip->i_itemp->ili_last_lsn does not store a log sequence number. What it stores is passed to it from the ->iop_committing method, which is called by xfs_log_commit_cil(). The value this passes to the iop_committing method is the CIL context sequence number that the item was committed to. As it turns out, xlog_cil_force_lsn() converts the sequence to an actual commit LSN for the related context and returns that to xfs_log_force_lsn(). xfs_log_force_lsn() overwrites it's "lsn" variable that contained a sequence with an actual LSN and then uses that to sync the iclogs. This caused me some confusion for a while, even though I originally wrote all this code a decade ago. ->iop_committing is only used by a couple of log item types, and only inode items use the sequence number it is passed. Let's clean up the API, CIL structures and inode log item to call it a sequence number, and make it clear that the high level code is using CIL sequence numbers and not on-disk LSNs for integrity synchronisation purposes. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Allison Henderson <allison.henderson@oracle.com> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-06-18 23:21:52 +08:00
xfs_csn_t seq;
xfs_ilock(ip, XFS_ILOCK_SHARED);
xfs: xfs_log_force_lsn isn't passed a LSN In doing an investigation into AIL push stalls, I was looking at the log force code to see if an async CIL push could be done instead. This lead me to xfs_log_force_lsn() and looking at how it works. xfs_log_force_lsn() is only called from inode synchronisation contexts such as fsync(), and it takes the ip->i_itemp->ili_last_lsn value as the LSN to sync the log to. This gets passed to xlog_cil_force_lsn() via xfs_log_force_lsn() to flush the CIL to the journal, and then used by xfs_log_force_lsn() to flush the iclogs to the journal. The problem is that ip->i_itemp->ili_last_lsn does not store a log sequence number. What it stores is passed to it from the ->iop_committing method, which is called by xfs_log_commit_cil(). The value this passes to the iop_committing method is the CIL context sequence number that the item was committed to. As it turns out, xlog_cil_force_lsn() converts the sequence to an actual commit LSN for the related context and returns that to xfs_log_force_lsn(). xfs_log_force_lsn() overwrites it's "lsn" variable that contained a sequence with an actual LSN and then uses that to sync the iclogs. This caused me some confusion for a while, even though I originally wrote all this code a decade ago. ->iop_committing is only used by a couple of log item types, and only inode items use the sequence number it is passed. Let's clean up the API, CIL structures and inode log item to call it a sequence number, and make it clear that the high level code is using CIL sequence numbers and not on-disk LSNs for integrity synchronisation purposes. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Allison Henderson <allison.henderson@oracle.com> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-06-18 23:21:52 +08:00
seq = xfs_fsync_seq(ip, datasync);
if (seq) {
error = xfs_log_force_seq(ip->i_mount, seq, XFS_LOG_SYNC,
log_flushed);
spin_lock(&ip->i_itemp->ili_lock);
ip->i_itemp->ili_fsync_fields = 0;
spin_unlock(&ip->i_itemp->ili_lock);
}
xfs_iunlock(ip, XFS_ILOCK_SHARED);
return error;
}
STATIC int
xfs_file_fsync(
struct file *file,
loff_t start,
loff_t end,
int datasync)
{
struct xfs_inode *ip = XFS_I(file->f_mapping->host);
struct xfs_mount *mp = ip->i_mount;
int error, err2;
int log_flushed = 0;
trace_xfs_file_fsync(ip);
error = file_write_and_wait_range(file, start, end);
if (error)
return error;
if (xfs_is_shutdown(mp))
return -EIO;
xfs_iflags_clear(ip, XFS_ITRUNCATED);
/*
* If we have an RT and/or log subvolume we need to make sure to flush
* the write cache the device used for file data first. This is to
* ensure newly written file data make it to disk before logging the new
* inode size in case of an extending write.
*/
if (XFS_IS_REALTIME_INODE(ip))
error = blkdev_issue_flush(mp->m_rtdev_targp->bt_bdev);
else if (mp->m_logdev_targp != mp->m_ddev_targp)
error = blkdev_issue_flush(mp->m_ddev_targp->bt_bdev);
/*
* Any inode that has dirty modifications in the log is pinned. The
* racy check here for a pinned inode will not catch modifications
* that happen concurrently to the fsync call, but fsync semantics
* only require to sync previously completed I/O.
*/
if (xfs_ipincount(ip)) {
err2 = xfs_fsync_flush_log(ip, datasync, &log_flushed);
if (err2 && !error)
error = err2;
}
/*
* If we only have a single device, and the log force about was
* a no-op we might have to flush the data device cache here.
* This can only happen for fdatasync/O_DSYNC if we were overwriting
* an already allocated file and thus do not have any metadata to
* commit.
*/
if (!log_flushed && !XFS_IS_REALTIME_INODE(ip) &&
mp->m_logdev_targp == mp->m_ddev_targp) {
err2 = blkdev_issue_flush(mp->m_ddev_targp->bt_bdev);
if (err2 && !error)
error = err2;
}
return error;
}
static int
xfs_ilock_iocb(
struct kiocb *iocb,
unsigned int lock_mode)
{
struct xfs_inode *ip = XFS_I(file_inode(iocb->ki_filp));
if (iocb->ki_flags & IOCB_NOWAIT) {
if (!xfs_ilock_nowait(ip, lock_mode))
return -EAGAIN;
} else {
xfs_ilock(ip, lock_mode);
}
return 0;
}
STATIC ssize_t
xfs_file_dio_read(
struct kiocb *iocb,
struct iov_iter *to)
{
struct xfs_inode *ip = XFS_I(file_inode(iocb->ki_filp));
ssize_t ret;
trace_xfs_file_direct_read(iocb, to);
if (!iov_iter_count(to))
return 0; /* skip atime */
file_accessed(iocb->ki_filp);
ret = xfs_ilock_iocb(iocb, XFS_IOLOCK_SHARED);
if (ret)
return ret;
ret = iomap_dio_rw(iocb, to, &xfs_read_iomap_ops, NULL, 0, NULL, 0);
xfs_iunlock(ip, XFS_IOLOCK_SHARED);
return ret;
}
static noinline ssize_t
xfs_file_dax_read(
struct kiocb *iocb,
struct iov_iter *to)
{
struct xfs_inode *ip = XFS_I(iocb->ki_filp->f_mapping->host);
ssize_t ret = 0;
trace_xfs_file_dax_read(iocb, to);
if (!iov_iter_count(to))
return 0; /* skip atime */
ret = xfs_ilock_iocb(iocb, XFS_IOLOCK_SHARED);
if (ret)
return ret;
ret = dax_iomap_rw(iocb, to, &xfs_read_iomap_ops);
xfs_iunlock(ip, XFS_IOLOCK_SHARED);
file_accessed(iocb->ki_filp);
return ret;
}
STATIC ssize_t
xfs_file_buffered_read(
struct kiocb *iocb,
struct iov_iter *to)
{
struct xfs_inode *ip = XFS_I(file_inode(iocb->ki_filp));
ssize_t ret;
trace_xfs_file_buffered_read(iocb, to);
ret = xfs_ilock_iocb(iocb, XFS_IOLOCK_SHARED);
if (ret)
return ret;
ret = generic_file_read_iter(iocb, to);
xfs_iunlock(ip, XFS_IOLOCK_SHARED);
return ret;
}
STATIC ssize_t
xfs_file_read_iter(
struct kiocb *iocb,
struct iov_iter *to)
{
struct inode *inode = file_inode(iocb->ki_filp);
struct xfs_mount *mp = XFS_I(inode)->i_mount;
ssize_t ret = 0;
XFS_STATS_INC(mp, xs_read_calls);
if (xfs_is_shutdown(mp))
return -EIO;
if (IS_DAX(inode))
ret = xfs_file_dax_read(iocb, to);
else if (iocb->ki_flags & IOCB_DIRECT)
ret = xfs_file_dio_read(iocb, to);
else
ret = xfs_file_buffered_read(iocb, to);
if (ret > 0)
XFS_STATS_ADD(mp, xs_read_bytes, ret);
return ret;
}
/*
* Common pre-write limit and setup checks.
*
* Called with the iolocked held either shared and exclusive according to
* @iolock, and returns with it held. Might upgrade the iolock to exclusive
* if called for a direct write beyond i_size.
*/
STATIC ssize_t
xfs_file_write_checks(
struct kiocb *iocb,
struct iov_iter *from,
unsigned int *iolock)
{
struct file *file = iocb->ki_filp;
struct inode *inode = file->f_mapping->host;
struct xfs_inode *ip = XFS_I(inode);
ssize_t error = 0;
size_t count = iov_iter_count(from);
xfs: always drain dio before extending aio write submission XFS supports and typically allows concurrent asynchronous direct I/O submission to a single file. One exception to the rule is that file extending dio writes that start beyond the current EOF (e.g., potentially create a hole at EOF) require exclusive I/O access to the file. This is because such writes must zero any pre-existing blocks beyond EOF that are exposed by virtue of now residing within EOF as a result of the write about to be submitted. Before EOF zeroing can occur, the current file i_size must be stabilized to avoid data corruption. In this scenario, XFS upgrades the iolock to exclude any further I/O submission, waits on in-flight I/O to complete to ensure i_size is up to date (i_size is updated on dio write completion) and restarts the various checks against the state of the file. The problem is that this protection sequence is triggered only when the iolock is currently held shared. While this is true for async dio in most cases, the caller may upgrade the lock in advance based on arbitrary circumstances with respect to EOF zeroing. For example, the iolock is always acquired exclusively if the start offset is not block aligned. This means that even though the iolock is already held exclusive for such I/Os, pending I/O is not drained and thus EOF zeroing can occur based on an unstable i_size. This problem has been reproduced as guest data corruption in virtual machines with file-backed qcow2 virtual disks hosted on an XFS filesystem. The virtual disks must be configured with aio=native mode and the must not be truncated out to the maximum file size (as some virt managers will do). Update xfs_file_aio_write_checks() to unconditionally drain in-flight dio before EOF zeroing can occur. Rather than trigger the wait based on iolock state, use a new flag and upgrade the iolock when necessary. Note that this results in a full restart of the inode checks even when the iolock was already held exclusive when technically it is only required to recheck i_size. This should be a rare enough occurrence that it is preferable to keep the code simple rather than create an alternate restart jump target. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Eric Sandeen <sandeen@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2015-10-12 13:02:05 +08:00
bool drained_dio = false;
loff_t isize;
xfs: don't serialise adjacent concurrent direct IO appending writes For append write workloads, extending the file requires a certain amount of exclusive locking to be done up front to ensure sanity in things like ensuring that we've zeroed any allocated regions between the old EOF and the start of the new IO. For single threads, this typically isn't a problem, and for large IOs we don't serialise enough for it to be a problem for two threads on really fast block devices. However for smaller IO and larger thread counts we have a problem. Take 4 concurrent sequential, single block sized and aligned IOs. After the first IO is submitted but before it completes, we end up with this state: IO 1 IO 2 IO 3 IO 4 +-------+-------+-------+-------+ ^ ^ | | | | | | | \- ip->i_new_size \- ip->i_size And the IO is done without exclusive locking because offset <= ip->i_size. When we submit IO 2, we see offset > ip->i_size, and grab the IO lock exclusive, because there is a chance we need to do EOF zeroing. However, there is already an IO in progress that avoids the need for IO zeroing because offset <= ip->i_new_size. hence we could avoid holding the IO lock exlcusive for this. Hence after submission of the second IO, we'd end up this state: IO 1 IO 2 IO 3 IO 4 +-------+-------+-------+-------+ ^ ^ | | | | | | | \- ip->i_new_size \- ip->i_size There is no need to grab the i_mutex of the IO lock in exclusive mode if we don't need to invalidate the page cache. Taking these locks on every direct IO effective serialises them as taking the IO lock in exclusive mode has to wait for all shared holders to drop the lock. That only happens when IO is complete, so effective it prevents dispatch of concurrent direct IO writes to the same inode. And so you can see that for the third concurrent IO, we'd avoid exclusive locking for the same reason we avoided the exclusive lock for the second IO. Fixing this is a bit more complex than that, because we need to hold a write-submission local value of ip->i_new_size to that clearing the value is only done if no other thread has updated it before our IO completes..... Signed-off-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
2011-08-25 15:17:02 +08:00
restart:
error = generic_write_checks(iocb, from);
if (error <= 0)
return error;
if (iocb->ki_flags & IOCB_NOWAIT) {
error = break_layout(inode, false);
if (error == -EWOULDBLOCK)
error = -EAGAIN;
} else {
error = xfs_break_layouts(inode, iolock, BREAK_WRITE);
}
if (error)
return error;
/*
* For changing security info in file_remove_privs() we need i_rwsem
* exclusively.
*/
if (*iolock == XFS_IOLOCK_SHARED && !IS_NOSEC(inode)) {
xfs_iunlock(ip, *iolock);
*iolock = XFS_IOLOCK_EXCL;
error = xfs_ilock_iocb(iocb, *iolock);
if (error) {
*iolock = 0;
return error;
}
goto restart;
}
/*
* If the offset is beyond the size of the file, we need to zero any
* blocks that fall between the existing EOF and the start of this
* write. If zeroing is needed and we are currently holding the iolock
* shared, we need to update it to exclusive which implies having to
* redo all checks before.
*
* We need to serialise against EOF updates that occur in IO completions
* here. We want to make sure that nobody is changing the size while we
* do this check until we have placed an IO barrier (i.e. hold the
* XFS_IOLOCK_EXCL) that prevents new IO from being dispatched. The
* spinlock effectively forms a memory barrier once we have the
* XFS_IOLOCK_EXCL so we are guaranteed to see the latest EOF value and
* hence be able to correctly determine if we need to run zeroing.
xfs: DIO write completion size updates race xfs_end_io_direct_write() can race with other IO completions when updating the in-core inode size. The IO completion processing is not serialised for direct IO - they are done either under the IOLOCK_SHARED for non-AIO DIO, and without any IOLOCK held at all during AIO DIO completion. Hence the non-atomic test-and-set update of the in-core inode size is racy and can result in the in-core inode size going backwards if the race if hit just right. If the inode size goes backwards, this can trigger the EOF zeroing code to run incorrectly on the next IO, which then will zero data that has successfully been written to disk by a previous DIO. To fix this bug, we need to serialise the test/set updates of the in-core inode size. This first patch introduces locking around the relevant updates and checks in the DIO path. Because we now have an ioend in xfs_end_io_direct_write(), we know exactly then we are doing an IO that requires an in-core EOF update, and we know that they are not running in interrupt context. As such, we do not need to use irqsave() spinlock variants to protect against interrupts while the lock is held. Hence we can use an existing spinlock in the inode to do this serialisation and so not need to grow the struct xfs_inode just to work around this problem. This patch does not address the test/set EOF update in generic_file_write_direct() for various reasons - that will be done as a followup with separate explanation. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2015-04-16 20:03:07 +08:00
*
* We can do an unlocked check here safely as IO completion can only
* extend EOF. Truncate is locked out at this point, so the EOF can
* not move backwards, only forwards. Hence we only need to take the
* slow path and spin locks when we are at or beyond the current EOF.
*/
if (iocb->ki_pos <= i_size_read(inode))
goto out;
xfs: DIO write completion size updates race xfs_end_io_direct_write() can race with other IO completions when updating the in-core inode size. The IO completion processing is not serialised for direct IO - they are done either under the IOLOCK_SHARED for non-AIO DIO, and without any IOLOCK held at all during AIO DIO completion. Hence the non-atomic test-and-set update of the in-core inode size is racy and can result in the in-core inode size going backwards if the race if hit just right. If the inode size goes backwards, this can trigger the EOF zeroing code to run incorrectly on the next IO, which then will zero data that has successfully been written to disk by a previous DIO. To fix this bug, we need to serialise the test/set updates of the in-core inode size. This first patch introduces locking around the relevant updates and checks in the DIO path. Because we now have an ioend in xfs_end_io_direct_write(), we know exactly then we are doing an IO that requires an in-core EOF update, and we know that they are not running in interrupt context. As such, we do not need to use irqsave() spinlock variants to protect against interrupts while the lock is held. Hence we can use an existing spinlock in the inode to do this serialisation and so not need to grow the struct xfs_inode just to work around this problem. This patch does not address the test/set EOF update in generic_file_write_direct() for various reasons - that will be done as a followup with separate explanation. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2015-04-16 20:03:07 +08:00
spin_lock(&ip->i_flags_lock);
isize = i_size_read(inode);
if (iocb->ki_pos > isize) {
xfs: DIO write completion size updates race xfs_end_io_direct_write() can race with other IO completions when updating the in-core inode size. The IO completion processing is not serialised for direct IO - they are done either under the IOLOCK_SHARED for non-AIO DIO, and without any IOLOCK held at all during AIO DIO completion. Hence the non-atomic test-and-set update of the in-core inode size is racy and can result in the in-core inode size going backwards if the race if hit just right. If the inode size goes backwards, this can trigger the EOF zeroing code to run incorrectly on the next IO, which then will zero data that has successfully been written to disk by a previous DIO. To fix this bug, we need to serialise the test/set updates of the in-core inode size. This first patch introduces locking around the relevant updates and checks in the DIO path. Because we now have an ioend in xfs_end_io_direct_write(), we know exactly then we are doing an IO that requires an in-core EOF update, and we know that they are not running in interrupt context. As such, we do not need to use irqsave() spinlock variants to protect against interrupts while the lock is held. Hence we can use an existing spinlock in the inode to do this serialisation and so not need to grow the struct xfs_inode just to work around this problem. This patch does not address the test/set EOF update in generic_file_write_direct() for various reasons - that will be done as a followup with separate explanation. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2015-04-16 20:03:07 +08:00
spin_unlock(&ip->i_flags_lock);
if (iocb->ki_flags & IOCB_NOWAIT)
return -EAGAIN;
xfs: always drain dio before extending aio write submission XFS supports and typically allows concurrent asynchronous direct I/O submission to a single file. One exception to the rule is that file extending dio writes that start beyond the current EOF (e.g., potentially create a hole at EOF) require exclusive I/O access to the file. This is because such writes must zero any pre-existing blocks beyond EOF that are exposed by virtue of now residing within EOF as a result of the write about to be submitted. Before EOF zeroing can occur, the current file i_size must be stabilized to avoid data corruption. In this scenario, XFS upgrades the iolock to exclude any further I/O submission, waits on in-flight I/O to complete to ensure i_size is up to date (i_size is updated on dio write completion) and restarts the various checks against the state of the file. The problem is that this protection sequence is triggered only when the iolock is currently held shared. While this is true for async dio in most cases, the caller may upgrade the lock in advance based on arbitrary circumstances with respect to EOF zeroing. For example, the iolock is always acquired exclusively if the start offset is not block aligned. This means that even though the iolock is already held exclusive for such I/Os, pending I/O is not drained and thus EOF zeroing can occur based on an unstable i_size. This problem has been reproduced as guest data corruption in virtual machines with file-backed qcow2 virtual disks hosted on an XFS filesystem. The virtual disks must be configured with aio=native mode and the must not be truncated out to the maximum file size (as some virt managers will do). Update xfs_file_aio_write_checks() to unconditionally drain in-flight dio before EOF zeroing can occur. Rather than trigger the wait based on iolock state, use a new flag and upgrade the iolock when necessary. Note that this results in a full restart of the inode checks even when the iolock was already held exclusive when technically it is only required to recheck i_size. This should be a rare enough occurrence that it is preferable to keep the code simple rather than create an alternate restart jump target. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Eric Sandeen <sandeen@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2015-10-12 13:02:05 +08:00
if (!drained_dio) {
if (*iolock == XFS_IOLOCK_SHARED) {
xfs_iunlock(ip, *iolock);
xfs: always drain dio before extending aio write submission XFS supports and typically allows concurrent asynchronous direct I/O submission to a single file. One exception to the rule is that file extending dio writes that start beyond the current EOF (e.g., potentially create a hole at EOF) require exclusive I/O access to the file. This is because such writes must zero any pre-existing blocks beyond EOF that are exposed by virtue of now residing within EOF as a result of the write about to be submitted. Before EOF zeroing can occur, the current file i_size must be stabilized to avoid data corruption. In this scenario, XFS upgrades the iolock to exclude any further I/O submission, waits on in-flight I/O to complete to ensure i_size is up to date (i_size is updated on dio write completion) and restarts the various checks against the state of the file. The problem is that this protection sequence is triggered only when the iolock is currently held shared. While this is true for async dio in most cases, the caller may upgrade the lock in advance based on arbitrary circumstances with respect to EOF zeroing. For example, the iolock is always acquired exclusively if the start offset is not block aligned. This means that even though the iolock is already held exclusive for such I/Os, pending I/O is not drained and thus EOF zeroing can occur based on an unstable i_size. This problem has been reproduced as guest data corruption in virtual machines with file-backed qcow2 virtual disks hosted on an XFS filesystem. The virtual disks must be configured with aio=native mode and the must not be truncated out to the maximum file size (as some virt managers will do). Update xfs_file_aio_write_checks() to unconditionally drain in-flight dio before EOF zeroing can occur. Rather than trigger the wait based on iolock state, use a new flag and upgrade the iolock when necessary. Note that this results in a full restart of the inode checks even when the iolock was already held exclusive when technically it is only required to recheck i_size. This should be a rare enough occurrence that it is preferable to keep the code simple rather than create an alternate restart jump target. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Eric Sandeen <sandeen@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2015-10-12 13:02:05 +08:00
*iolock = XFS_IOLOCK_EXCL;
xfs_ilock(ip, *iolock);
xfs: always drain dio before extending aio write submission XFS supports and typically allows concurrent asynchronous direct I/O submission to a single file. One exception to the rule is that file extending dio writes that start beyond the current EOF (e.g., potentially create a hole at EOF) require exclusive I/O access to the file. This is because such writes must zero any pre-existing blocks beyond EOF that are exposed by virtue of now residing within EOF as a result of the write about to be submitted. Before EOF zeroing can occur, the current file i_size must be stabilized to avoid data corruption. In this scenario, XFS upgrades the iolock to exclude any further I/O submission, waits on in-flight I/O to complete to ensure i_size is up to date (i_size is updated on dio write completion) and restarts the various checks against the state of the file. The problem is that this protection sequence is triggered only when the iolock is currently held shared. While this is true for async dio in most cases, the caller may upgrade the lock in advance based on arbitrary circumstances with respect to EOF zeroing. For example, the iolock is always acquired exclusively if the start offset is not block aligned. This means that even though the iolock is already held exclusive for such I/Os, pending I/O is not drained and thus EOF zeroing can occur based on an unstable i_size. This problem has been reproduced as guest data corruption in virtual machines with file-backed qcow2 virtual disks hosted on an XFS filesystem. The virtual disks must be configured with aio=native mode and the must not be truncated out to the maximum file size (as some virt managers will do). Update xfs_file_aio_write_checks() to unconditionally drain in-flight dio before EOF zeroing can occur. Rather than trigger the wait based on iolock state, use a new flag and upgrade the iolock when necessary. Note that this results in a full restart of the inode checks even when the iolock was already held exclusive when technically it is only required to recheck i_size. This should be a rare enough occurrence that it is preferable to keep the code simple rather than create an alternate restart jump target. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Eric Sandeen <sandeen@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2015-10-12 13:02:05 +08:00
iov_iter_reexpand(from, count);
}
/*
* We now have an IO submission barrier in place, but
* AIO can do EOF updates during IO completion and hence
* we now need to wait for all of them to drain. Non-AIO
* DIO will have drained before we are given the
* XFS_IOLOCK_EXCL, and so for most cases this wait is a
* no-op.
*/
inode_dio_wait(inode);
xfs: always drain dio before extending aio write submission XFS supports and typically allows concurrent asynchronous direct I/O submission to a single file. One exception to the rule is that file extending dio writes that start beyond the current EOF (e.g., potentially create a hole at EOF) require exclusive I/O access to the file. This is because such writes must zero any pre-existing blocks beyond EOF that are exposed by virtue of now residing within EOF as a result of the write about to be submitted. Before EOF zeroing can occur, the current file i_size must be stabilized to avoid data corruption. In this scenario, XFS upgrades the iolock to exclude any further I/O submission, waits on in-flight I/O to complete to ensure i_size is up to date (i_size is updated on dio write completion) and restarts the various checks against the state of the file. The problem is that this protection sequence is triggered only when the iolock is currently held shared. While this is true for async dio in most cases, the caller may upgrade the lock in advance based on arbitrary circumstances with respect to EOF zeroing. For example, the iolock is always acquired exclusively if the start offset is not block aligned. This means that even though the iolock is already held exclusive for such I/Os, pending I/O is not drained and thus EOF zeroing can occur based on an unstable i_size. This problem has been reproduced as guest data corruption in virtual machines with file-backed qcow2 virtual disks hosted on an XFS filesystem. The virtual disks must be configured with aio=native mode and the must not be truncated out to the maximum file size (as some virt managers will do). Update xfs_file_aio_write_checks() to unconditionally drain in-flight dio before EOF zeroing can occur. Rather than trigger the wait based on iolock state, use a new flag and upgrade the iolock when necessary. Note that this results in a full restart of the inode checks even when the iolock was already held exclusive when technically it is only required to recheck i_size. This should be a rare enough occurrence that it is preferable to keep the code simple rather than create an alternate restart jump target. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Eric Sandeen <sandeen@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2015-10-12 13:02:05 +08:00
drained_dio = true;
xfs: don't serialise adjacent concurrent direct IO appending writes For append write workloads, extending the file requires a certain amount of exclusive locking to be done up front to ensure sanity in things like ensuring that we've zeroed any allocated regions between the old EOF and the start of the new IO. For single threads, this typically isn't a problem, and for large IOs we don't serialise enough for it to be a problem for two threads on really fast block devices. However for smaller IO and larger thread counts we have a problem. Take 4 concurrent sequential, single block sized and aligned IOs. After the first IO is submitted but before it completes, we end up with this state: IO 1 IO 2 IO 3 IO 4 +-------+-------+-------+-------+ ^ ^ | | | | | | | \- ip->i_new_size \- ip->i_size And the IO is done without exclusive locking because offset <= ip->i_size. When we submit IO 2, we see offset > ip->i_size, and grab the IO lock exclusive, because there is a chance we need to do EOF zeroing. However, there is already an IO in progress that avoids the need for IO zeroing because offset <= ip->i_new_size. hence we could avoid holding the IO lock exlcusive for this. Hence after submission of the second IO, we'd end up this state: IO 1 IO 2 IO 3 IO 4 +-------+-------+-------+-------+ ^ ^ | | | | | | | \- ip->i_new_size \- ip->i_size There is no need to grab the i_mutex of the IO lock in exclusive mode if we don't need to invalidate the page cache. Taking these locks on every direct IO effective serialises them as taking the IO lock in exclusive mode has to wait for all shared holders to drop the lock. That only happens when IO is complete, so effective it prevents dispatch of concurrent direct IO writes to the same inode. And so you can see that for the third concurrent IO, we'd avoid exclusive locking for the same reason we avoided the exclusive lock for the second IO. Fixing this is a bit more complex than that, because we need to hold a write-submission local value of ip->i_new_size to that clearing the value is only done if no other thread has updated it before our IO completes..... Signed-off-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
2011-08-25 15:17:02 +08:00
goto restart;
}
trace_xfs_zero_eof(ip, isize, iocb->ki_pos - isize);
error = xfs_zero_range(ip, isize, iocb->ki_pos - isize, NULL);
if (error)
return error;
xfs: DIO write completion size updates race xfs_end_io_direct_write() can race with other IO completions when updating the in-core inode size. The IO completion processing is not serialised for direct IO - they are done either under the IOLOCK_SHARED for non-AIO DIO, and without any IOLOCK held at all during AIO DIO completion. Hence the non-atomic test-and-set update of the in-core inode size is racy and can result in the in-core inode size going backwards if the race if hit just right. If the inode size goes backwards, this can trigger the EOF zeroing code to run incorrectly on the next IO, which then will zero data that has successfully been written to disk by a previous DIO. To fix this bug, we need to serialise the test/set updates of the in-core inode size. This first patch introduces locking around the relevant updates and checks in the DIO path. Because we now have an ioend in xfs_end_io_direct_write(), we know exactly then we are doing an IO that requires an in-core EOF update, and we know that they are not running in interrupt context. As such, we do not need to use irqsave() spinlock variants to protect against interrupts while the lock is held. Hence we can use an existing spinlock in the inode to do this serialisation and so not need to grow the struct xfs_inode just to work around this problem. This patch does not address the test/set EOF update in generic_file_write_direct() for various reasons - that will be done as a followup with separate explanation. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2015-04-16 20:03:07 +08:00
} else
spin_unlock(&ip->i_flags_lock);
out:
return kiocb_modified(iocb);
}
static int
xfs_dio_write_end_io(
struct kiocb *iocb,
ssize_t size,
int error,
unsigned flags)
{
struct inode *inode = file_inode(iocb->ki_filp);
struct xfs_inode *ip = XFS_I(inode);
loff_t offset = iocb->ki_pos;
unsigned int nofs_flag;
trace_xfs_end_io_direct_write(ip, offset, size);
if (xfs_is_shutdown(ip->i_mount))
return -EIO;
if (error)
return error;
if (!size)
return 0;
/*
* Capture amount written on completion as we can't reliably account
* for it on submission.
*/
XFS_STATS_ADD(ip->i_mount, xs_write_bytes, size);
/*
* We can allocate memory here while doing writeback on behalf of
* memory reclaim. To avoid memory allocation deadlocks set the
* task-wide nofs context for the following operations.
*/
nofs_flag = memalloc_nofs_save();
if (flags & IOMAP_DIO_COW) {
error = xfs_reflink_end_cow(ip, offset, size);
if (error)
goto out;
}
/*
* Unwritten conversion updates the in-core isize after extent
* conversion but before updating the on-disk size. Updating isize any
* earlier allows a racing dio read to find unwritten extents before
* they are converted.
*/
if (flags & IOMAP_DIO_UNWRITTEN) {
error = xfs_iomap_write_unwritten(ip, offset, size, true);
goto out;
}
/*
* We need to update the in-core inode size here so that we don't end up
* with the on-disk inode size being outside the in-core inode size. We
* have no other method of updating EOF for AIO, so always do it here
* if necessary.
*
* We need to lock the test/set EOF update as we can be racing with
* other IO completions here to update the EOF. Failing to serialise
* here can result in EOF moving backwards and Bad Things Happen when
* that occurs.
*
* As IO completion only ever extends EOF, we can do an unlocked check
* here to avoid taking the spinlock. If we land within the current EOF,
* then we do not need to do an extending update at all, and we don't
* need to take the lock to check this. If we race with an update moving
* EOF, then we'll either still be beyond EOF and need to take the lock,
* or we'll be within EOF and we don't need to take it at all.
*/
if (offset + size <= i_size_read(inode))
goto out;
spin_lock(&ip->i_flags_lock);
if (offset + size > i_size_read(inode)) {
i_size_write(inode, offset + size);
spin_unlock(&ip->i_flags_lock);
error = xfs_setfilesize(ip, offset, size);
} else {
spin_unlock(&ip->i_flags_lock);
}
out:
memalloc_nofs_restore(nofs_flag);
return error;
}
static const struct iomap_dio_ops xfs_dio_write_ops = {
.end_io = xfs_dio_write_end_io,
};
/*
* Handle block aligned direct I/O writes
*/
static noinline ssize_t
xfs_file_dio_write_aligned(
struct xfs_inode *ip,
struct kiocb *iocb,
struct iov_iter *from)
{
unsigned int iolock = XFS_IOLOCK_SHARED;
ssize_t ret;
ret = xfs_ilock_iocb(iocb, iolock);
if (ret)
return ret;
ret = xfs_file_write_checks(iocb, from, &iolock);
if (ret)
goto out_unlock;
xfs: don't serialise adjacent concurrent direct IO appending writes For append write workloads, extending the file requires a certain amount of exclusive locking to be done up front to ensure sanity in things like ensuring that we've zeroed any allocated regions between the old EOF and the start of the new IO. For single threads, this typically isn't a problem, and for large IOs we don't serialise enough for it to be a problem for two threads on really fast block devices. However for smaller IO and larger thread counts we have a problem. Take 4 concurrent sequential, single block sized and aligned IOs. After the first IO is submitted but before it completes, we end up with this state: IO 1 IO 2 IO 3 IO 4 +-------+-------+-------+-------+ ^ ^ | | | | | | | \- ip->i_new_size \- ip->i_size And the IO is done without exclusive locking because offset <= ip->i_size. When we submit IO 2, we see offset > ip->i_size, and grab the IO lock exclusive, because there is a chance we need to do EOF zeroing. However, there is already an IO in progress that avoids the need for IO zeroing because offset <= ip->i_new_size. hence we could avoid holding the IO lock exlcusive for this. Hence after submission of the second IO, we'd end up this state: IO 1 IO 2 IO 3 IO 4 +-------+-------+-------+-------+ ^ ^ | | | | | | | \- ip->i_new_size \- ip->i_size There is no need to grab the i_mutex of the IO lock in exclusive mode if we don't need to invalidate the page cache. Taking these locks on every direct IO effective serialises them as taking the IO lock in exclusive mode has to wait for all shared holders to drop the lock. That only happens when IO is complete, so effective it prevents dispatch of concurrent direct IO writes to the same inode. And so you can see that for the third concurrent IO, we'd avoid exclusive locking for the same reason we avoided the exclusive lock for the second IO. Fixing this is a bit more complex than that, because we need to hold a write-submission local value of ip->i_new_size to that clearing the value is only done if no other thread has updated it before our IO completes..... Signed-off-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
2011-08-25 15:17:02 +08:00
/*
* We don't need to hold the IOLOCK exclusively across the IO, so demote
* the iolock back to shared if we had to take the exclusive lock in
* xfs_file_write_checks() for other reasons.
xfs: don't serialise adjacent concurrent direct IO appending writes For append write workloads, extending the file requires a certain amount of exclusive locking to be done up front to ensure sanity in things like ensuring that we've zeroed any allocated regions between the old EOF and the start of the new IO. For single threads, this typically isn't a problem, and for large IOs we don't serialise enough for it to be a problem for two threads on really fast block devices. However for smaller IO and larger thread counts we have a problem. Take 4 concurrent sequential, single block sized and aligned IOs. After the first IO is submitted but before it completes, we end up with this state: IO 1 IO 2 IO 3 IO 4 +-------+-------+-------+-------+ ^ ^ | | | | | | | \- ip->i_new_size \- ip->i_size And the IO is done without exclusive locking because offset <= ip->i_size. When we submit IO 2, we see offset > ip->i_size, and grab the IO lock exclusive, because there is a chance we need to do EOF zeroing. However, there is already an IO in progress that avoids the need for IO zeroing because offset <= ip->i_new_size. hence we could avoid holding the IO lock exlcusive for this. Hence after submission of the second IO, we'd end up this state: IO 1 IO 2 IO 3 IO 4 +-------+-------+-------+-------+ ^ ^ | | | | | | | \- ip->i_new_size \- ip->i_size There is no need to grab the i_mutex of the IO lock in exclusive mode if we don't need to invalidate the page cache. Taking these locks on every direct IO effective serialises them as taking the IO lock in exclusive mode has to wait for all shared holders to drop the lock. That only happens when IO is complete, so effective it prevents dispatch of concurrent direct IO writes to the same inode. And so you can see that for the third concurrent IO, we'd avoid exclusive locking for the same reason we avoided the exclusive lock for the second IO. Fixing this is a bit more complex than that, because we need to hold a write-submission local value of ip->i_new_size to that clearing the value is only done if no other thread has updated it before our IO completes..... Signed-off-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
2011-08-25 15:17:02 +08:00
*/
if (iolock == XFS_IOLOCK_EXCL) {
xfs_ilock_demote(ip, XFS_IOLOCK_EXCL);
iolock = XFS_IOLOCK_SHARED;
}
trace_xfs_file_direct_write(iocb, from);
ret = iomap_dio_rw(iocb, from, &xfs_direct_write_iomap_ops,
&xfs_dio_write_ops, 0, NULL, 0);
out_unlock:
if (iolock)
xfs_iunlock(ip, iolock);
return ret;
}
/*
* Handle block unaligned direct I/O writes
*
* In most cases direct I/O writes will be done holding IOLOCK_SHARED, allowing
* them to be done in parallel with reads and other direct I/O writes. However,
* if the I/O is not aligned to filesystem blocks, the direct I/O layer may need
* to do sub-block zeroing and that requires serialisation against other direct
* I/O to the same block. In this case we need to serialise the submission of
* the unaligned I/O so that we don't get racing block zeroing in the dio layer.
xfs: reduce exclusive locking on unaligned dio Attempt shared locking for unaligned DIO, but only if the the underlying extent is already allocated and in written state. On failure, retry with the existing exclusive locking. Test case is fio randrw of 512 byte IOs using AIO and an iodepth of 32 IOs. Vanilla: READ: bw=4560KiB/s (4670kB/s), 4560KiB/s-4560KiB/s (4670kB/s-4670kB/s), io=134MiB (140MB), run=30001-30001msec WRITE: bw=4567KiB/s (4676kB/s), 4567KiB/s-4567KiB/s (4676kB/s-4676kB/s), io=134MiB (140MB), run=30001-30001msec Patched: READ: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1127MiB (1182MB), run=30002-30002msec WRITE: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1128MiB (1183MB), run=30002-30002msec That's an improvement from ~18k IOPS to a ~150k IOPS, which is about the IOPS limit of the VM block device setup I'm testing on. 4kB block IO comparison: READ: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8868MiB (9299MB), run=30002-30002msec WRITE: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8878MiB (9309MB), run=30002-30002msec Which is ~150k IOPS, same as what the test gets for sub-block AIO+DIO writes with this patch. Signed-off-by: Dave Chinner <dchinner@redhat.com> [hch: rebased, split unaligned from nowait] Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-01-24 02:06:31 +08:00
* In the case where sub-block zeroing is not required, we can do concurrent
* sub-block dios to the same block successfully.
*
xfs: reduce exclusive locking on unaligned dio Attempt shared locking for unaligned DIO, but only if the the underlying extent is already allocated and in written state. On failure, retry with the existing exclusive locking. Test case is fio randrw of 512 byte IOs using AIO and an iodepth of 32 IOs. Vanilla: READ: bw=4560KiB/s (4670kB/s), 4560KiB/s-4560KiB/s (4670kB/s-4670kB/s), io=134MiB (140MB), run=30001-30001msec WRITE: bw=4567KiB/s (4676kB/s), 4567KiB/s-4567KiB/s (4676kB/s-4676kB/s), io=134MiB (140MB), run=30001-30001msec Patched: READ: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1127MiB (1182MB), run=30002-30002msec WRITE: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1128MiB (1183MB), run=30002-30002msec That's an improvement from ~18k IOPS to a ~150k IOPS, which is about the IOPS limit of the VM block device setup I'm testing on. 4kB block IO comparison: READ: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8868MiB (9299MB), run=30002-30002msec WRITE: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8878MiB (9309MB), run=30002-30002msec Which is ~150k IOPS, same as what the test gets for sub-block AIO+DIO writes with this patch. Signed-off-by: Dave Chinner <dchinner@redhat.com> [hch: rebased, split unaligned from nowait] Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-01-24 02:06:31 +08:00
* Optimistically submit the I/O using the shared lock first, but use the
* IOMAP_DIO_OVERWRITE_ONLY flag to tell the lower layers to return -EAGAIN
* if block allocation or partial block zeroing would be required. In that case
* we try again with the exclusive lock.
*/
static noinline ssize_t
xfs_file_dio_write_unaligned(
struct xfs_inode *ip,
struct kiocb *iocb,
struct iov_iter *from)
{
xfs: reduce exclusive locking on unaligned dio Attempt shared locking for unaligned DIO, but only if the the underlying extent is already allocated and in written state. On failure, retry with the existing exclusive locking. Test case is fio randrw of 512 byte IOs using AIO and an iodepth of 32 IOs. Vanilla: READ: bw=4560KiB/s (4670kB/s), 4560KiB/s-4560KiB/s (4670kB/s-4670kB/s), io=134MiB (140MB), run=30001-30001msec WRITE: bw=4567KiB/s (4676kB/s), 4567KiB/s-4567KiB/s (4676kB/s-4676kB/s), io=134MiB (140MB), run=30001-30001msec Patched: READ: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1127MiB (1182MB), run=30002-30002msec WRITE: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1128MiB (1183MB), run=30002-30002msec That's an improvement from ~18k IOPS to a ~150k IOPS, which is about the IOPS limit of the VM block device setup I'm testing on. 4kB block IO comparison: READ: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8868MiB (9299MB), run=30002-30002msec WRITE: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8878MiB (9309MB), run=30002-30002msec Which is ~150k IOPS, same as what the test gets for sub-block AIO+DIO writes with this patch. Signed-off-by: Dave Chinner <dchinner@redhat.com> [hch: rebased, split unaligned from nowait] Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-01-24 02:06:31 +08:00
size_t isize = i_size_read(VFS_I(ip));
size_t count = iov_iter_count(from);
unsigned int iolock = XFS_IOLOCK_SHARED;
xfs: reduce exclusive locking on unaligned dio Attempt shared locking for unaligned DIO, but only if the the underlying extent is already allocated and in written state. On failure, retry with the existing exclusive locking. Test case is fio randrw of 512 byte IOs using AIO and an iodepth of 32 IOs. Vanilla: READ: bw=4560KiB/s (4670kB/s), 4560KiB/s-4560KiB/s (4670kB/s-4670kB/s), io=134MiB (140MB), run=30001-30001msec WRITE: bw=4567KiB/s (4676kB/s), 4567KiB/s-4567KiB/s (4676kB/s-4676kB/s), io=134MiB (140MB), run=30001-30001msec Patched: READ: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1127MiB (1182MB), run=30002-30002msec WRITE: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1128MiB (1183MB), run=30002-30002msec That's an improvement from ~18k IOPS to a ~150k IOPS, which is about the IOPS limit of the VM block device setup I'm testing on. 4kB block IO comparison: READ: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8868MiB (9299MB), run=30002-30002msec WRITE: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8878MiB (9309MB), run=30002-30002msec Which is ~150k IOPS, same as what the test gets for sub-block AIO+DIO writes with this patch. Signed-off-by: Dave Chinner <dchinner@redhat.com> [hch: rebased, split unaligned from nowait] Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-01-24 02:06:31 +08:00
unsigned int flags = IOMAP_DIO_OVERWRITE_ONLY;
ssize_t ret;
xfs: reduce exclusive locking on unaligned dio Attempt shared locking for unaligned DIO, but only if the the underlying extent is already allocated and in written state. On failure, retry with the existing exclusive locking. Test case is fio randrw of 512 byte IOs using AIO and an iodepth of 32 IOs. Vanilla: READ: bw=4560KiB/s (4670kB/s), 4560KiB/s-4560KiB/s (4670kB/s-4670kB/s), io=134MiB (140MB), run=30001-30001msec WRITE: bw=4567KiB/s (4676kB/s), 4567KiB/s-4567KiB/s (4676kB/s-4676kB/s), io=134MiB (140MB), run=30001-30001msec Patched: READ: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1127MiB (1182MB), run=30002-30002msec WRITE: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1128MiB (1183MB), run=30002-30002msec That's an improvement from ~18k IOPS to a ~150k IOPS, which is about the IOPS limit of the VM block device setup I'm testing on. 4kB block IO comparison: READ: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8868MiB (9299MB), run=30002-30002msec WRITE: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8878MiB (9309MB), run=30002-30002msec Which is ~150k IOPS, same as what the test gets for sub-block AIO+DIO writes with this patch. Signed-off-by: Dave Chinner <dchinner@redhat.com> [hch: rebased, split unaligned from nowait] Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-01-24 02:06:31 +08:00
/*
* Extending writes need exclusivity because of the sub-block zeroing
* that the DIO code always does for partial tail blocks beyond EOF, so
* don't even bother trying the fast path in this case.
*/
if (iocb->ki_pos > isize || iocb->ki_pos + count >= isize) {
if (iocb->ki_flags & IOCB_NOWAIT)
return -EAGAIN;
retry_exclusive:
xfs: reduce exclusive locking on unaligned dio Attempt shared locking for unaligned DIO, but only if the the underlying extent is already allocated and in written state. On failure, retry with the existing exclusive locking. Test case is fio randrw of 512 byte IOs using AIO and an iodepth of 32 IOs. Vanilla: READ: bw=4560KiB/s (4670kB/s), 4560KiB/s-4560KiB/s (4670kB/s-4670kB/s), io=134MiB (140MB), run=30001-30001msec WRITE: bw=4567KiB/s (4676kB/s), 4567KiB/s-4567KiB/s (4676kB/s-4676kB/s), io=134MiB (140MB), run=30001-30001msec Patched: READ: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1127MiB (1182MB), run=30002-30002msec WRITE: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1128MiB (1183MB), run=30002-30002msec That's an improvement from ~18k IOPS to a ~150k IOPS, which is about the IOPS limit of the VM block device setup I'm testing on. 4kB block IO comparison: READ: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8868MiB (9299MB), run=30002-30002msec WRITE: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8878MiB (9309MB), run=30002-30002msec Which is ~150k IOPS, same as what the test gets for sub-block AIO+DIO writes with this patch. Signed-off-by: Dave Chinner <dchinner@redhat.com> [hch: rebased, split unaligned from nowait] Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-01-24 02:06:31 +08:00
iolock = XFS_IOLOCK_EXCL;
flags = IOMAP_DIO_FORCE_WAIT;
}
ret = xfs_ilock_iocb(iocb, iolock);
if (ret)
return ret;
/*
* We can't properly handle unaligned direct I/O to reflink files yet,
* as we can't unshare a partial block.
*/
if (xfs_is_cow_inode(ip)) {
trace_xfs_reflink_bounce_dio_write(iocb, from);
ret = -ENOTBLK;
goto out_unlock;
}
ret = xfs_file_write_checks(iocb, from, &iolock);
if (ret)
goto out_unlock;
/*
xfs: reduce exclusive locking on unaligned dio Attempt shared locking for unaligned DIO, but only if the the underlying extent is already allocated and in written state. On failure, retry with the existing exclusive locking. Test case is fio randrw of 512 byte IOs using AIO and an iodepth of 32 IOs. Vanilla: READ: bw=4560KiB/s (4670kB/s), 4560KiB/s-4560KiB/s (4670kB/s-4670kB/s), io=134MiB (140MB), run=30001-30001msec WRITE: bw=4567KiB/s (4676kB/s), 4567KiB/s-4567KiB/s (4676kB/s-4676kB/s), io=134MiB (140MB), run=30001-30001msec Patched: READ: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1127MiB (1182MB), run=30002-30002msec WRITE: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1128MiB (1183MB), run=30002-30002msec That's an improvement from ~18k IOPS to a ~150k IOPS, which is about the IOPS limit of the VM block device setup I'm testing on. 4kB block IO comparison: READ: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8868MiB (9299MB), run=30002-30002msec WRITE: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8878MiB (9309MB), run=30002-30002msec Which is ~150k IOPS, same as what the test gets for sub-block AIO+DIO writes with this patch. Signed-off-by: Dave Chinner <dchinner@redhat.com> [hch: rebased, split unaligned from nowait] Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-01-24 02:06:31 +08:00
* If we are doing exclusive unaligned I/O, this must be the only I/O
* in-flight. Otherwise we risk data corruption due to unwritten extent
* conversions from the AIO end_io handler. Wait for all other I/O to
* drain first.
*/
xfs: reduce exclusive locking on unaligned dio Attempt shared locking for unaligned DIO, but only if the the underlying extent is already allocated and in written state. On failure, retry with the existing exclusive locking. Test case is fio randrw of 512 byte IOs using AIO and an iodepth of 32 IOs. Vanilla: READ: bw=4560KiB/s (4670kB/s), 4560KiB/s-4560KiB/s (4670kB/s-4670kB/s), io=134MiB (140MB), run=30001-30001msec WRITE: bw=4567KiB/s (4676kB/s), 4567KiB/s-4567KiB/s (4676kB/s-4676kB/s), io=134MiB (140MB), run=30001-30001msec Patched: READ: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1127MiB (1182MB), run=30002-30002msec WRITE: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1128MiB (1183MB), run=30002-30002msec That's an improvement from ~18k IOPS to a ~150k IOPS, which is about the IOPS limit of the VM block device setup I'm testing on. 4kB block IO comparison: READ: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8868MiB (9299MB), run=30002-30002msec WRITE: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8878MiB (9309MB), run=30002-30002msec Which is ~150k IOPS, same as what the test gets for sub-block AIO+DIO writes with this patch. Signed-off-by: Dave Chinner <dchinner@redhat.com> [hch: rebased, split unaligned from nowait] Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-01-24 02:06:31 +08:00
if (flags & IOMAP_DIO_FORCE_WAIT)
inode_dio_wait(VFS_I(ip));
trace_xfs_file_direct_write(iocb, from);
ret = iomap_dio_rw(iocb, from, &xfs_direct_write_iomap_ops,
&xfs_dio_write_ops, flags, NULL, 0);
xfs: reduce exclusive locking on unaligned dio Attempt shared locking for unaligned DIO, but only if the the underlying extent is already allocated and in written state. On failure, retry with the existing exclusive locking. Test case is fio randrw of 512 byte IOs using AIO and an iodepth of 32 IOs. Vanilla: READ: bw=4560KiB/s (4670kB/s), 4560KiB/s-4560KiB/s (4670kB/s-4670kB/s), io=134MiB (140MB), run=30001-30001msec WRITE: bw=4567KiB/s (4676kB/s), 4567KiB/s-4567KiB/s (4676kB/s-4676kB/s), io=134MiB (140MB), run=30001-30001msec Patched: READ: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1127MiB (1182MB), run=30002-30002msec WRITE: bw=37.6MiB/s (39.4MB/s), 37.6MiB/s-37.6MiB/s (39.4MB/s-39.4MB/s), io=1128MiB (1183MB), run=30002-30002msec That's an improvement from ~18k IOPS to a ~150k IOPS, which is about the IOPS limit of the VM block device setup I'm testing on. 4kB block IO comparison: READ: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8868MiB (9299MB), run=30002-30002msec WRITE: bw=296MiB/s (310MB/s), 296MiB/s-296MiB/s (310MB/s-310MB/s), io=8878MiB (9309MB), run=30002-30002msec Which is ~150k IOPS, same as what the test gets for sub-block AIO+DIO writes with this patch. Signed-off-by: Dave Chinner <dchinner@redhat.com> [hch: rebased, split unaligned from nowait] Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-01-24 02:06:31 +08:00
/*
* Retry unaligned I/O with exclusive blocking semantics if the DIO
* layer rejected it for mapping or locking reasons. If we are doing
* nonblocking user I/O, propagate the error.
*/
if (ret == -EAGAIN && !(iocb->ki_flags & IOCB_NOWAIT)) {
ASSERT(flags & IOMAP_DIO_OVERWRITE_ONLY);
xfs_iunlock(ip, iolock);
goto retry_exclusive;
}
out_unlock:
if (iolock)
xfs_iunlock(ip, iolock);
return ret;
}
static ssize_t
xfs_file_dio_write(
struct kiocb *iocb,
struct iov_iter *from)
{
struct xfs_inode *ip = XFS_I(file_inode(iocb->ki_filp));
struct xfs_buftarg *target = xfs_inode_buftarg(ip);
size_t count = iov_iter_count(from);
/* direct I/O must be aligned to device logical sector size */
if ((iocb->ki_pos | count) & target->bt_logical_sectormask)
return -EINVAL;
if ((iocb->ki_pos | count) & ip->i_mount->m_blockmask)
return xfs_file_dio_write_unaligned(ip, iocb, from);
return xfs_file_dio_write_aligned(ip, iocb, from);
}
static noinline ssize_t
xfs_file_dax_write(
struct kiocb *iocb,
struct iov_iter *from)
{
struct inode *inode = iocb->ki_filp->f_mapping->host;
struct xfs_inode *ip = XFS_I(inode);
unsigned int iolock = XFS_IOLOCK_EXCL;
ssize_t ret, error = 0;
loff_t pos;
ret = xfs_ilock_iocb(iocb, iolock);
if (ret)
return ret;
ret = xfs_file_write_checks(iocb, from, &iolock);
if (ret)
goto out;
pos = iocb->ki_pos;
trace_xfs_file_dax_write(iocb, from);
ret = dax_iomap_rw(iocb, from, &xfs_dax_write_iomap_ops);
if (ret > 0 && iocb->ki_pos > i_size_read(inode)) {
i_size_write(inode, iocb->ki_pos);
error = xfs_setfilesize(ip, pos, ret);
}
out:
if (iolock)
xfs_iunlock(ip, iolock);
if (error)
return error;
if (ret > 0) {
XFS_STATS_ADD(ip->i_mount, xs_write_bytes, ret);
/* Handle various SYNC-type writes */
ret = generic_write_sync(iocb, ret);
}
return ret;
}
STATIC ssize_t
xfs_file_buffered_write(
struct kiocb *iocb,
struct iov_iter *from)
{
struct inode *inode = iocb->ki_filp->f_mapping->host;
struct xfs_inode *ip = XFS_I(inode);
ssize_t ret;
bool cleared_space = false;
unsigned int iolock;
xfs: sync eofblocks scans under iolock are livelock prone The xfs_eofblocks.eof_scan_owner field is an internal field to facilitate invoking eofb scans from the kernel while under the iolock. This is necessary because the eofb scan acquires the iolock of each inode. Synchronous scans are invoked on certain buffered write failures while under iolock. In such cases, the scan owner indicates that the context for the scan already owns the particular iolock and prevents a double lock deadlock. eofblocks scans while under iolock are still livelock prone in the event of multiple parallel scans, however. If multiple buffered writes to different inodes fail and invoke eofblocks scans at the same time, each scan avoids a deadlock with its own inode by virtue of the eof_scan_owner field, but will never be able to acquire the iolock of the inode from the parallel scan. Because the low free space scans are invoked with SYNC_WAIT, the scan will not return until it has processed every tagged inode and thus both scans will spin indefinitely on the iolock being held across the opposite scan. This problem can be reproduced reliably by generic/224 on systems with higher cpu counts (x16). To avoid this problem, simplify the semantics of eofblocks scans to never invoke a scan while under iolock. This means that the buffered write context must drop the iolock before the scan. It must reacquire the lock before the write retry and also repeat the initial write checks, as the original state might no longer be valid once the iolock was dropped. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
2017-01-28 15:22:56 +08:00
write_retry:
iolock = XFS_IOLOCK_EXCL;
ret = xfs_ilock_iocb(iocb, iolock);
if (ret)
return ret;
ret = xfs_file_write_checks(iocb, from, &iolock);
if (ret)
goto out;
/* We can write back this queue in page reclaim */
current->backing_dev_info = inode_to_bdi(inode);
trace_xfs_file_buffered_write(iocb, from);
ret = iomap_file_buffered_write(iocb, from,
&xfs_buffered_write_iomap_ops);
if (likely(ret >= 0))
iocb->ki_pos += ret;
xfs: run an eofblocks scan on ENOSPC/EDQUOT From: Brian Foster <bfoster@redhat.com> Speculative preallocation and and the associated throttling metrics assume we're working with large files on large filesystems. Users have reported inefficiencies in these mechanisms when we happen to be dealing with large files on smaller filesystems. This can occur because while prealloc throttling is aggressive under low free space conditions, it is not active until we reach 5% free space or less. For example, a 40GB filesystem has enough space for several files large enough to have multi-GB preallocations at any given time. If those files are slow growing, they might reserve preallocation for long periods of time as well as avoid the background scanner due to frequent modification. If a new file is written under these conditions, said file has no access to this already reserved space and premature ENOSPC is imminent. To handle this scenario, modify the buffered write ENOSPC handling and retry sequence to invoke an eofblocks scan. In the smaller filesystem scenario, the eofblocks scan resets the usage of preallocation such that when the 5% free space threshold is met, throttling effectively takes over to provide fair and efficient preallocation until legitimate ENOSPC. The eofblocks scan is selective based on the nature of the failure. For example, an EDQUOT failure in a particular quota will use a filtered scan for that quota. Because we don't know which quota might have caused an allocation failure at any given time, we include each applicable quota determined to be under low free space conditions in the scan. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2014-07-24 17:49:28 +08:00
/*
xfs: run an eofblocks scan on ENOSPC/EDQUOT From: Brian Foster <bfoster@redhat.com> Speculative preallocation and and the associated throttling metrics assume we're working with large files on large filesystems. Users have reported inefficiencies in these mechanisms when we happen to be dealing with large files on smaller filesystems. This can occur because while prealloc throttling is aggressive under low free space conditions, it is not active until we reach 5% free space or less. For example, a 40GB filesystem has enough space for several files large enough to have multi-GB preallocations at any given time. If those files are slow growing, they might reserve preallocation for long periods of time as well as avoid the background scanner due to frequent modification. If a new file is written under these conditions, said file has no access to this already reserved space and premature ENOSPC is imminent. To handle this scenario, modify the buffered write ENOSPC handling and retry sequence to invoke an eofblocks scan. In the smaller filesystem scenario, the eofblocks scan resets the usage of preallocation such that when the 5% free space threshold is met, throttling effectively takes over to provide fair and efficient preallocation until legitimate ENOSPC. The eofblocks scan is selective based on the nature of the failure. For example, an EDQUOT failure in a particular quota will use a filtered scan for that quota. Because we don't know which quota might have caused an allocation failure at any given time, we include each applicable quota determined to be under low free space conditions in the scan. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2014-07-24 17:49:28 +08:00
* If we hit a space limit, try to free up some lingering preallocated
* space before returning an error. In the case of ENOSPC, first try to
* write back all dirty inodes to free up some of the excess reserved
* metadata space. This reduces the chances that the eofblocks scan
* waits on dirty mappings. Since xfs_flush_inodes() is serialized, this
* also behaves as a filter to prevent too many eofblocks scans from
* running at the same time. Use a synchronous scan to increase the
* effectiveness of the scan.
*/
if (ret == -EDQUOT && !cleared_space) {
xfs: sync eofblocks scans under iolock are livelock prone The xfs_eofblocks.eof_scan_owner field is an internal field to facilitate invoking eofb scans from the kernel while under the iolock. This is necessary because the eofb scan acquires the iolock of each inode. Synchronous scans are invoked on certain buffered write failures while under iolock. In such cases, the scan owner indicates that the context for the scan already owns the particular iolock and prevents a double lock deadlock. eofblocks scans while under iolock are still livelock prone in the event of multiple parallel scans, however. If multiple buffered writes to different inodes fail and invoke eofblocks scans at the same time, each scan avoids a deadlock with its own inode by virtue of the eof_scan_owner field, but will never be able to acquire the iolock of the inode from the parallel scan. Because the low free space scans are invoked with SYNC_WAIT, the scan will not return until it has processed every tagged inode and thus both scans will spin indefinitely on the iolock being held across the opposite scan. This problem can be reproduced reliably by generic/224 on systems with higher cpu counts (x16). To avoid this problem, simplify the semantics of eofblocks scans to never invoke a scan while under iolock. This means that the buffered write context must drop the iolock before the scan. It must reacquire the lock before the write retry and also repeat the initial write checks, as the original state might no longer be valid once the iolock was dropped. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
2017-01-28 15:22:56 +08:00
xfs_iunlock(ip, iolock);
xfs_blockgc_free_quota(ip, XFS_ICWALK_FLAG_SYNC);
cleared_space = true;
goto write_retry;
} else if (ret == -ENOSPC && !cleared_space) {
struct xfs_icwalk icw = {0};
xfs: run an eofblocks scan on ENOSPC/EDQUOT From: Brian Foster <bfoster@redhat.com> Speculative preallocation and and the associated throttling metrics assume we're working with large files on large filesystems. Users have reported inefficiencies in these mechanisms when we happen to be dealing with large files on smaller filesystems. This can occur because while prealloc throttling is aggressive under low free space conditions, it is not active until we reach 5% free space or less. For example, a 40GB filesystem has enough space for several files large enough to have multi-GB preallocations at any given time. If those files are slow growing, they might reserve preallocation for long periods of time as well as avoid the background scanner due to frequent modification. If a new file is written under these conditions, said file has no access to this already reserved space and premature ENOSPC is imminent. To handle this scenario, modify the buffered write ENOSPC handling and retry sequence to invoke an eofblocks scan. In the smaller filesystem scenario, the eofblocks scan resets the usage of preallocation such that when the 5% free space threshold is met, throttling effectively takes over to provide fair and efficient preallocation until legitimate ENOSPC. The eofblocks scan is selective based on the nature of the failure. For example, an EDQUOT failure in a particular quota will use a filtered scan for that quota. Because we don't know which quota might have caused an allocation failure at any given time, we include each applicable quota determined to be under low free space conditions in the scan. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2014-07-24 17:49:28 +08:00
cleared_space = true;
xfs: xfs_sync_data is redundant. We don't do any data writeback from XFS any more - the VFS is completely responsible for that, including for freeze. We can replace the remaining caller with a VFS level function that achieves the same thing, but without conflicting with current writeback work. This means we can remove the flush_work and xfs_flush_inodes() - the VFS functionality completely replaces the internal flush queue for doing this writeback work in a separate context to avoid stack overruns. This does have one complication - it cannot be called with page locks held. Hence move the flushing of delalloc space when ENOSPC occurs back up into xfs_file_aio_buffered_write when we don't hold any locks that will stall writeback. Unfortunately, writeback_inodes_sb_if_idle() is not sufficient to trigger delalloc conversion fast enough to prevent spurious ENOSPC whent here are hundreds of writers, thousands of small files and GBs of free RAM. Hence we need to use sync_sb_inodes() to block callers while we wait for writeback like the previous xfs_flush_inodes implementation did. That means we have to hold the s_umount lock here, but because this call can nest inside i_mutex (the parent directory in the create case, held by the VFS), we have to use down_read_trylock() to avoid potential deadlocks. In practice, this trylock will succeed on almost every attempt as unmount/remount type operations are exceedingly rare. Note: we always need to pass a count of zero to generic_file_buffered_write() as the previously written byte count. We only do this by accident before this patch by the virtue of ret always being zero when there are no errors. Make this explicit rather than needing to specifically zero ret in the ENOSPC retry case. Signed-off-by: Dave Chinner <dchinner@redhat.com> Tested-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Ben Myers <bpm@sgi.com>
2012-10-08 18:56:04 +08:00
xfs_flush_inodes(ip->i_mount);
xfs: sync eofblocks scans under iolock are livelock prone The xfs_eofblocks.eof_scan_owner field is an internal field to facilitate invoking eofb scans from the kernel while under the iolock. This is necessary because the eofb scan acquires the iolock of each inode. Synchronous scans are invoked on certain buffered write failures while under iolock. In such cases, the scan owner indicates that the context for the scan already owns the particular iolock and prevents a double lock deadlock. eofblocks scans while under iolock are still livelock prone in the event of multiple parallel scans, however. If multiple buffered writes to different inodes fail and invoke eofblocks scans at the same time, each scan avoids a deadlock with its own inode by virtue of the eof_scan_owner field, but will never be able to acquire the iolock of the inode from the parallel scan. Because the low free space scans are invoked with SYNC_WAIT, the scan will not return until it has processed every tagged inode and thus both scans will spin indefinitely on the iolock being held across the opposite scan. This problem can be reproduced reliably by generic/224 on systems with higher cpu counts (x16). To avoid this problem, simplify the semantics of eofblocks scans to never invoke a scan while under iolock. This means that the buffered write context must drop the iolock before the scan. It must reacquire the lock before the write retry and also repeat the initial write checks, as the original state might no longer be valid once the iolock was dropped. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
2017-01-28 15:22:56 +08:00
xfs_iunlock(ip, iolock);
icw.icw_flags = XFS_ICWALK_FLAG_SYNC;
xfs_blockgc_free_space(ip->i_mount, &icw);
xfs: xfs_sync_data is redundant. We don't do any data writeback from XFS any more - the VFS is completely responsible for that, including for freeze. We can replace the remaining caller with a VFS level function that achieves the same thing, but without conflicting with current writeback work. This means we can remove the flush_work and xfs_flush_inodes() - the VFS functionality completely replaces the internal flush queue for doing this writeback work in a separate context to avoid stack overruns. This does have one complication - it cannot be called with page locks held. Hence move the flushing of delalloc space when ENOSPC occurs back up into xfs_file_aio_buffered_write when we don't hold any locks that will stall writeback. Unfortunately, writeback_inodes_sb_if_idle() is not sufficient to trigger delalloc conversion fast enough to prevent spurious ENOSPC whent here are hundreds of writers, thousands of small files and GBs of free RAM. Hence we need to use sync_sb_inodes() to block callers while we wait for writeback like the previous xfs_flush_inodes implementation did. That means we have to hold the s_umount lock here, but because this call can nest inside i_mutex (the parent directory in the create case, held by the VFS), we have to use down_read_trylock() to avoid potential deadlocks. In practice, this trylock will succeed on almost every attempt as unmount/remount type operations are exceedingly rare. Note: we always need to pass a count of zero to generic_file_buffered_write() as the previously written byte count. We only do this by accident before this patch by the virtue of ret always being zero when there are no errors. Make this explicit rather than needing to specifically zero ret in the ENOSPC retry case. Signed-off-by: Dave Chinner <dchinner@redhat.com> Tested-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Ben Myers <bpm@sgi.com>
2012-10-08 18:56:04 +08:00
goto write_retry;
}
current->backing_dev_info = NULL;
out:
xfs: sync eofblocks scans under iolock are livelock prone The xfs_eofblocks.eof_scan_owner field is an internal field to facilitate invoking eofb scans from the kernel while under the iolock. This is necessary because the eofb scan acquires the iolock of each inode. Synchronous scans are invoked on certain buffered write failures while under iolock. In such cases, the scan owner indicates that the context for the scan already owns the particular iolock and prevents a double lock deadlock. eofblocks scans while under iolock are still livelock prone in the event of multiple parallel scans, however. If multiple buffered writes to different inodes fail and invoke eofblocks scans at the same time, each scan avoids a deadlock with its own inode by virtue of the eof_scan_owner field, but will never be able to acquire the iolock of the inode from the parallel scan. Because the low free space scans are invoked with SYNC_WAIT, the scan will not return until it has processed every tagged inode and thus both scans will spin indefinitely on the iolock being held across the opposite scan. This problem can be reproduced reliably by generic/224 on systems with higher cpu counts (x16). To avoid this problem, simplify the semantics of eofblocks scans to never invoke a scan while under iolock. This means that the buffered write context must drop the iolock before the scan. It must reacquire the lock before the write retry and also repeat the initial write checks, as the original state might no longer be valid once the iolock was dropped. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
2017-01-28 15:22:56 +08:00
if (iolock)
xfs_iunlock(ip, iolock);
if (ret > 0) {
XFS_STATS_ADD(ip->i_mount, xs_write_bytes, ret);
/* Handle various SYNC-type writes */
ret = generic_write_sync(iocb, ret);
}
return ret;
}
STATIC ssize_t
xfs_file_write_iter(
struct kiocb *iocb,
struct iov_iter *from)
{
struct inode *inode = iocb->ki_filp->f_mapping->host;
struct xfs_inode *ip = XFS_I(inode);
ssize_t ret;
size_t ocount = iov_iter_count(from);
XFS_STATS_INC(ip->i_mount, xs_write_calls);
if (ocount == 0)
return 0;
if (xfs_is_shutdown(ip->i_mount))
return -EIO;
if (IS_DAX(inode))
return xfs_file_dax_write(iocb, from);
if (iocb->ki_flags & IOCB_DIRECT) {
/*
* Allow a directio write to fall back to a buffered
* write *only* in the case that we're doing a reflink
* CoW. In all other directio scenarios we do not
* allow an operation to fall back to buffered mode.
*/
ret = xfs_file_dio_write(iocb, from);
if (ret != -ENOTBLK)
return ret;
}
return xfs_file_buffered_write(iocb, from);
}
xfs, dax: introduce xfs_break_dax_layouts() xfs_break_dax_layouts(), similar to xfs_break_leased_layouts(), scans for busy / pinned dax pages and waits for those pages to go idle before any potential extent unmap operation. dax_layout_busy_page() handles synchronizing against new page-busy events (get_user_pages). It invalidates all mappings to trigger the get_user_pages slow path which will eventually block on the xfs inode lock held in XFS_MMAPLOCK_EXCL mode. If dax_layout_busy_page() finds a busy page it returns it for xfs to wait for the page-idle event that will fire when the page reference count reaches 1 (recall ZONE_DEVICE pages are idle at count 1, see generic_dax_pagefree()). While waiting, the XFS_MMAPLOCK_EXCL lock is dropped in order to not deadlock the process that might be trying to elevate the page count of more pages before arranging for any of them to go idle. I.e. the typical case of submitting I/O is that iov_iter_get_pages() elevates the reference count of all pages in the I/O before starting I/O on the first page. The process of elevating the reference count of all pages involved in an I/O may cause faults that need to take XFS_MMAPLOCK_EXCL. Although XFS_MMAPLOCK_EXCL is dropped while waiting, XFS_IOLOCK_EXCL is held while sleeping. We need this to prevent starvation of the truncate path as continuous submission of direct-I/O could starve the truncate path indefinitely if the lock is dropped. Cc: Dave Chinner <david@fromorbit.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Reported-by: Jan Kara <jack@suse.cz> Reviewed-by: Jan Kara <jack@suse.cz> Reviewed-by: Christoph Hellwig <hch@lst.de> Acked-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-05-10 06:47:49 +08:00
static void
xfs_wait_dax_page(
struct inode *inode)
xfs, dax: introduce xfs_break_dax_layouts() xfs_break_dax_layouts(), similar to xfs_break_leased_layouts(), scans for busy / pinned dax pages and waits for those pages to go idle before any potential extent unmap operation. dax_layout_busy_page() handles synchronizing against new page-busy events (get_user_pages). It invalidates all mappings to trigger the get_user_pages slow path which will eventually block on the xfs inode lock held in XFS_MMAPLOCK_EXCL mode. If dax_layout_busy_page() finds a busy page it returns it for xfs to wait for the page-idle event that will fire when the page reference count reaches 1 (recall ZONE_DEVICE pages are idle at count 1, see generic_dax_pagefree()). While waiting, the XFS_MMAPLOCK_EXCL lock is dropped in order to not deadlock the process that might be trying to elevate the page count of more pages before arranging for any of them to go idle. I.e. the typical case of submitting I/O is that iov_iter_get_pages() elevates the reference count of all pages in the I/O before starting I/O on the first page. The process of elevating the reference count of all pages involved in an I/O may cause faults that need to take XFS_MMAPLOCK_EXCL. Although XFS_MMAPLOCK_EXCL is dropped while waiting, XFS_IOLOCK_EXCL is held while sleeping. We need this to prevent starvation of the truncate path as continuous submission of direct-I/O could starve the truncate path indefinitely if the lock is dropped. Cc: Dave Chinner <david@fromorbit.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Reported-by: Jan Kara <jack@suse.cz> Reviewed-by: Jan Kara <jack@suse.cz> Reviewed-by: Christoph Hellwig <hch@lst.de> Acked-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-05-10 06:47:49 +08:00
{
struct xfs_inode *ip = XFS_I(inode);
xfs_iunlock(ip, XFS_MMAPLOCK_EXCL);
schedule();
xfs_ilock(ip, XFS_MMAPLOCK_EXCL);
}
int
xfs, dax: introduce xfs_break_dax_layouts() xfs_break_dax_layouts(), similar to xfs_break_leased_layouts(), scans for busy / pinned dax pages and waits for those pages to go idle before any potential extent unmap operation. dax_layout_busy_page() handles synchronizing against new page-busy events (get_user_pages). It invalidates all mappings to trigger the get_user_pages slow path which will eventually block on the xfs inode lock held in XFS_MMAPLOCK_EXCL mode. If dax_layout_busy_page() finds a busy page it returns it for xfs to wait for the page-idle event that will fire when the page reference count reaches 1 (recall ZONE_DEVICE pages are idle at count 1, see generic_dax_pagefree()). While waiting, the XFS_MMAPLOCK_EXCL lock is dropped in order to not deadlock the process that might be trying to elevate the page count of more pages before arranging for any of them to go idle. I.e. the typical case of submitting I/O is that iov_iter_get_pages() elevates the reference count of all pages in the I/O before starting I/O on the first page. The process of elevating the reference count of all pages involved in an I/O may cause faults that need to take XFS_MMAPLOCK_EXCL. Although XFS_MMAPLOCK_EXCL is dropped while waiting, XFS_IOLOCK_EXCL is held while sleeping. We need this to prevent starvation of the truncate path as continuous submission of direct-I/O could starve the truncate path indefinitely if the lock is dropped. Cc: Dave Chinner <david@fromorbit.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Reported-by: Jan Kara <jack@suse.cz> Reviewed-by: Jan Kara <jack@suse.cz> Reviewed-by: Christoph Hellwig <hch@lst.de> Acked-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-05-10 06:47:49 +08:00
xfs_break_dax_layouts(
struct inode *inode,
bool *retry)
xfs, dax: introduce xfs_break_dax_layouts() xfs_break_dax_layouts(), similar to xfs_break_leased_layouts(), scans for busy / pinned dax pages and waits for those pages to go idle before any potential extent unmap operation. dax_layout_busy_page() handles synchronizing against new page-busy events (get_user_pages). It invalidates all mappings to trigger the get_user_pages slow path which will eventually block on the xfs inode lock held in XFS_MMAPLOCK_EXCL mode. If dax_layout_busy_page() finds a busy page it returns it for xfs to wait for the page-idle event that will fire when the page reference count reaches 1 (recall ZONE_DEVICE pages are idle at count 1, see generic_dax_pagefree()). While waiting, the XFS_MMAPLOCK_EXCL lock is dropped in order to not deadlock the process that might be trying to elevate the page count of more pages before arranging for any of them to go idle. I.e. the typical case of submitting I/O is that iov_iter_get_pages() elevates the reference count of all pages in the I/O before starting I/O on the first page. The process of elevating the reference count of all pages involved in an I/O may cause faults that need to take XFS_MMAPLOCK_EXCL. Although XFS_MMAPLOCK_EXCL is dropped while waiting, XFS_IOLOCK_EXCL is held while sleeping. We need this to prevent starvation of the truncate path as continuous submission of direct-I/O could starve the truncate path indefinitely if the lock is dropped. Cc: Dave Chinner <david@fromorbit.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Reported-by: Jan Kara <jack@suse.cz> Reviewed-by: Jan Kara <jack@suse.cz> Reviewed-by: Christoph Hellwig <hch@lst.de> Acked-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-05-10 06:47:49 +08:00
{
struct page *page;
ASSERT(xfs_isilocked(XFS_I(inode), XFS_MMAPLOCK_EXCL));
page = dax_layout_busy_page(inode->i_mapping);
if (!page)
return 0;
*retry = true;
xfs, dax: introduce xfs_break_dax_layouts() xfs_break_dax_layouts(), similar to xfs_break_leased_layouts(), scans for busy / pinned dax pages and waits for those pages to go idle before any potential extent unmap operation. dax_layout_busy_page() handles synchronizing against new page-busy events (get_user_pages). It invalidates all mappings to trigger the get_user_pages slow path which will eventually block on the xfs inode lock held in XFS_MMAPLOCK_EXCL mode. If dax_layout_busy_page() finds a busy page it returns it for xfs to wait for the page-idle event that will fire when the page reference count reaches 1 (recall ZONE_DEVICE pages are idle at count 1, see generic_dax_pagefree()). While waiting, the XFS_MMAPLOCK_EXCL lock is dropped in order to not deadlock the process that might be trying to elevate the page count of more pages before arranging for any of them to go idle. I.e. the typical case of submitting I/O is that iov_iter_get_pages() elevates the reference count of all pages in the I/O before starting I/O on the first page. The process of elevating the reference count of all pages involved in an I/O may cause faults that need to take XFS_MMAPLOCK_EXCL. Although XFS_MMAPLOCK_EXCL is dropped while waiting, XFS_IOLOCK_EXCL is held while sleeping. We need this to prevent starvation of the truncate path as continuous submission of direct-I/O could starve the truncate path indefinitely if the lock is dropped. Cc: Dave Chinner <david@fromorbit.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Reported-by: Jan Kara <jack@suse.cz> Reviewed-by: Jan Kara <jack@suse.cz> Reviewed-by: Christoph Hellwig <hch@lst.de> Acked-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-05-10 06:47:49 +08:00
return ___wait_var_event(&page->_refcount,
atomic_read(&page->_refcount) == 1, TASK_INTERRUPTIBLE,
0, 0, xfs_wait_dax_page(inode));
xfs, dax: introduce xfs_break_dax_layouts() xfs_break_dax_layouts(), similar to xfs_break_leased_layouts(), scans for busy / pinned dax pages and waits for those pages to go idle before any potential extent unmap operation. dax_layout_busy_page() handles synchronizing against new page-busy events (get_user_pages). It invalidates all mappings to trigger the get_user_pages slow path which will eventually block on the xfs inode lock held in XFS_MMAPLOCK_EXCL mode. If dax_layout_busy_page() finds a busy page it returns it for xfs to wait for the page-idle event that will fire when the page reference count reaches 1 (recall ZONE_DEVICE pages are idle at count 1, see generic_dax_pagefree()). While waiting, the XFS_MMAPLOCK_EXCL lock is dropped in order to not deadlock the process that might be trying to elevate the page count of more pages before arranging for any of them to go idle. I.e. the typical case of submitting I/O is that iov_iter_get_pages() elevates the reference count of all pages in the I/O before starting I/O on the first page. The process of elevating the reference count of all pages involved in an I/O may cause faults that need to take XFS_MMAPLOCK_EXCL. Although XFS_MMAPLOCK_EXCL is dropped while waiting, XFS_IOLOCK_EXCL is held while sleeping. We need this to prevent starvation of the truncate path as continuous submission of direct-I/O could starve the truncate path indefinitely if the lock is dropped. Cc: Dave Chinner <david@fromorbit.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Reported-by: Jan Kara <jack@suse.cz> Reviewed-by: Jan Kara <jack@suse.cz> Reviewed-by: Christoph Hellwig <hch@lst.de> Acked-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-05-10 06:47:49 +08:00
}
xfs: prepare xfs_break_layouts() for another layout type When xfs is operating as the back-end of a pNFS block server, it prevents collisions between local and remote operations by requiring a lease to be held for remotely accessed blocks. Local filesystem operations break those leases before writing or mutating the extent map of the file. A similar mechanism is needed to prevent operations on pinned dax mappings, like device-DMA, from colliding with extent unmap operations. BREAK_WRITE and BREAK_UNMAP are introduced as two distinct levels of layout breaking. Layouts are broken in the BREAK_WRITE case to ensure that layout-holders do not collide with local writes. Additionally, layouts are broken in the BREAK_UNMAP case to make sure the layout-holder has a consistent view of the file's extent map. While BREAK_WRITE breaks can be satisfied be recalling FL_LAYOUT leases, BREAK_UNMAP breaks additionally require waiting for busy dax-pages to go idle while holding XFS_MMAPLOCK_EXCL. After this refactoring xfs_break_layouts() becomes the entry point for coordinating both types of breaks. Finally, xfs_break_leased_layouts() becomes just the BREAK_WRITE handler. Note that the unlock tracking is needed in a follow on change. That will coordinate retrying either break handler until both successfully test for a lease break while maintaining the lock state. Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Cc: "Darrick J. Wong" <darrick.wong@oracle.com> Reported-by: Dave Chinner <david@fromorbit.com> Reported-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-03-21 05:42:38 +08:00
int
xfs_break_layouts(
struct inode *inode,
uint *iolock,
enum layout_break_reason reason)
{
bool retry;
xfs, dax: introduce xfs_break_dax_layouts() xfs_break_dax_layouts(), similar to xfs_break_leased_layouts(), scans for busy / pinned dax pages and waits for those pages to go idle before any potential extent unmap operation. dax_layout_busy_page() handles synchronizing against new page-busy events (get_user_pages). It invalidates all mappings to trigger the get_user_pages slow path which will eventually block on the xfs inode lock held in XFS_MMAPLOCK_EXCL mode. If dax_layout_busy_page() finds a busy page it returns it for xfs to wait for the page-idle event that will fire when the page reference count reaches 1 (recall ZONE_DEVICE pages are idle at count 1, see generic_dax_pagefree()). While waiting, the XFS_MMAPLOCK_EXCL lock is dropped in order to not deadlock the process that might be trying to elevate the page count of more pages before arranging for any of them to go idle. I.e. the typical case of submitting I/O is that iov_iter_get_pages() elevates the reference count of all pages in the I/O before starting I/O on the first page. The process of elevating the reference count of all pages involved in an I/O may cause faults that need to take XFS_MMAPLOCK_EXCL. Although XFS_MMAPLOCK_EXCL is dropped while waiting, XFS_IOLOCK_EXCL is held while sleeping. We need this to prevent starvation of the truncate path as continuous submission of direct-I/O could starve the truncate path indefinitely if the lock is dropped. Cc: Dave Chinner <david@fromorbit.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Reported-by: Jan Kara <jack@suse.cz> Reviewed-by: Jan Kara <jack@suse.cz> Reviewed-by: Christoph Hellwig <hch@lst.de> Acked-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-05-10 06:47:49 +08:00
int error;
xfs: prepare xfs_break_layouts() for another layout type When xfs is operating as the back-end of a pNFS block server, it prevents collisions between local and remote operations by requiring a lease to be held for remotely accessed blocks. Local filesystem operations break those leases before writing or mutating the extent map of the file. A similar mechanism is needed to prevent operations on pinned dax mappings, like device-DMA, from colliding with extent unmap operations. BREAK_WRITE and BREAK_UNMAP are introduced as two distinct levels of layout breaking. Layouts are broken in the BREAK_WRITE case to ensure that layout-holders do not collide with local writes. Additionally, layouts are broken in the BREAK_UNMAP case to make sure the layout-holder has a consistent view of the file's extent map. While BREAK_WRITE breaks can be satisfied be recalling FL_LAYOUT leases, BREAK_UNMAP breaks additionally require waiting for busy dax-pages to go idle while holding XFS_MMAPLOCK_EXCL. After this refactoring xfs_break_layouts() becomes the entry point for coordinating both types of breaks. Finally, xfs_break_leased_layouts() becomes just the BREAK_WRITE handler. Note that the unlock tracking is needed in a follow on change. That will coordinate retrying either break handler until both successfully test for a lease break while maintaining the lock state. Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Cc: "Darrick J. Wong" <darrick.wong@oracle.com> Reported-by: Dave Chinner <david@fromorbit.com> Reported-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-03-21 05:42:38 +08:00
ASSERT(xfs_isilocked(XFS_I(inode), XFS_IOLOCK_SHARED|XFS_IOLOCK_EXCL));
xfs, dax: introduce xfs_break_dax_layouts() xfs_break_dax_layouts(), similar to xfs_break_leased_layouts(), scans for busy / pinned dax pages and waits for those pages to go idle before any potential extent unmap operation. dax_layout_busy_page() handles synchronizing against new page-busy events (get_user_pages). It invalidates all mappings to trigger the get_user_pages slow path which will eventually block on the xfs inode lock held in XFS_MMAPLOCK_EXCL mode. If dax_layout_busy_page() finds a busy page it returns it for xfs to wait for the page-idle event that will fire when the page reference count reaches 1 (recall ZONE_DEVICE pages are idle at count 1, see generic_dax_pagefree()). While waiting, the XFS_MMAPLOCK_EXCL lock is dropped in order to not deadlock the process that might be trying to elevate the page count of more pages before arranging for any of them to go idle. I.e. the typical case of submitting I/O is that iov_iter_get_pages() elevates the reference count of all pages in the I/O before starting I/O on the first page. The process of elevating the reference count of all pages involved in an I/O may cause faults that need to take XFS_MMAPLOCK_EXCL. Although XFS_MMAPLOCK_EXCL is dropped while waiting, XFS_IOLOCK_EXCL is held while sleeping. We need this to prevent starvation of the truncate path as continuous submission of direct-I/O could starve the truncate path indefinitely if the lock is dropped. Cc: Dave Chinner <david@fromorbit.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Reported-by: Jan Kara <jack@suse.cz> Reviewed-by: Jan Kara <jack@suse.cz> Reviewed-by: Christoph Hellwig <hch@lst.de> Acked-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-05-10 06:47:49 +08:00
do {
retry = false;
switch (reason) {
case BREAK_UNMAP:
error = xfs_break_dax_layouts(inode, &retry);
xfs, dax: introduce xfs_break_dax_layouts() xfs_break_dax_layouts(), similar to xfs_break_leased_layouts(), scans for busy / pinned dax pages and waits for those pages to go idle before any potential extent unmap operation. dax_layout_busy_page() handles synchronizing against new page-busy events (get_user_pages). It invalidates all mappings to trigger the get_user_pages slow path which will eventually block on the xfs inode lock held in XFS_MMAPLOCK_EXCL mode. If dax_layout_busy_page() finds a busy page it returns it for xfs to wait for the page-idle event that will fire when the page reference count reaches 1 (recall ZONE_DEVICE pages are idle at count 1, see generic_dax_pagefree()). While waiting, the XFS_MMAPLOCK_EXCL lock is dropped in order to not deadlock the process that might be trying to elevate the page count of more pages before arranging for any of them to go idle. I.e. the typical case of submitting I/O is that iov_iter_get_pages() elevates the reference count of all pages in the I/O before starting I/O on the first page. The process of elevating the reference count of all pages involved in an I/O may cause faults that need to take XFS_MMAPLOCK_EXCL. Although XFS_MMAPLOCK_EXCL is dropped while waiting, XFS_IOLOCK_EXCL is held while sleeping. We need this to prevent starvation of the truncate path as continuous submission of direct-I/O could starve the truncate path indefinitely if the lock is dropped. Cc: Dave Chinner <david@fromorbit.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Reported-by: Jan Kara <jack@suse.cz> Reviewed-by: Jan Kara <jack@suse.cz> Reviewed-by: Christoph Hellwig <hch@lst.de> Acked-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-05-10 06:47:49 +08:00
if (error || retry)
break;
xfs: Fix fall-through warnings for Clang In preparation to enable -Wimplicit-fallthrough for Clang, fix the following warnings by replacing /* fall through */ comments, and its variants, with the new pseudo-keyword macro fallthrough: fs/xfs/libxfs/xfs_alloc.c:3167:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/libxfs/xfs_da_btree.c:286:3: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/libxfs/xfs_ag_resv.c:346:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/libxfs/xfs_ag_resv.c:388:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/xfs_bmap_util.c:246:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/xfs_export.c:88:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/xfs_export.c:96:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/xfs_file.c:867:3: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/xfs_ioctl.c:562:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/xfs_ioctl.c:1548:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/xfs_iomap.c:1040:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/xfs_inode.c:852:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/xfs_log.c:2627:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/xfs_trans_buf.c:298:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/scrub/bmap.c:275:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/scrub/btree.c:48:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/scrub/common.c:85:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/scrub/common.c:138:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/scrub/common.c:698:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/scrub/dabtree.c:51:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/scrub/repair.c:951:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] fs/xfs/scrub/agheader.c:89:2: warning: unannotated fall-through between switch labels [-Wimplicit-fallthrough] Notice that Clang doesn't recognize /* fall through */ comments as implicit fall-through markings, so in order to globally enable -Wimplicit-fallthrough for Clang, these comments need to be replaced with fallthrough; in the whole codebase. Link: https://github.com/KSPP/linux/issues/115 Signed-off-by: Gustavo A. R. Silva <gustavoars@kernel.org>
2021-04-21 06:54:36 +08:00
fallthrough;
xfs, dax: introduce xfs_break_dax_layouts() xfs_break_dax_layouts(), similar to xfs_break_leased_layouts(), scans for busy / pinned dax pages and waits for those pages to go idle before any potential extent unmap operation. dax_layout_busy_page() handles synchronizing against new page-busy events (get_user_pages). It invalidates all mappings to trigger the get_user_pages slow path which will eventually block on the xfs inode lock held in XFS_MMAPLOCK_EXCL mode. If dax_layout_busy_page() finds a busy page it returns it for xfs to wait for the page-idle event that will fire when the page reference count reaches 1 (recall ZONE_DEVICE pages are idle at count 1, see generic_dax_pagefree()). While waiting, the XFS_MMAPLOCK_EXCL lock is dropped in order to not deadlock the process that might be trying to elevate the page count of more pages before arranging for any of them to go idle. I.e. the typical case of submitting I/O is that iov_iter_get_pages() elevates the reference count of all pages in the I/O before starting I/O on the first page. The process of elevating the reference count of all pages involved in an I/O may cause faults that need to take XFS_MMAPLOCK_EXCL. Although XFS_MMAPLOCK_EXCL is dropped while waiting, XFS_IOLOCK_EXCL is held while sleeping. We need this to prevent starvation of the truncate path as continuous submission of direct-I/O could starve the truncate path indefinitely if the lock is dropped. Cc: Dave Chinner <david@fromorbit.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Reported-by: Jan Kara <jack@suse.cz> Reviewed-by: Jan Kara <jack@suse.cz> Reviewed-by: Christoph Hellwig <hch@lst.de> Acked-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-05-10 06:47:49 +08:00
case BREAK_WRITE:
error = xfs_break_leased_layouts(inode, iolock, &retry);
break;
default:
WARN_ON_ONCE(1);
error = -EINVAL;
}
} while (error == 0 && retry);
return error;
xfs: prepare xfs_break_layouts() for another layout type When xfs is operating as the back-end of a pNFS block server, it prevents collisions between local and remote operations by requiring a lease to be held for remotely accessed blocks. Local filesystem operations break those leases before writing or mutating the extent map of the file. A similar mechanism is needed to prevent operations on pinned dax mappings, like device-DMA, from colliding with extent unmap operations. BREAK_WRITE and BREAK_UNMAP are introduced as two distinct levels of layout breaking. Layouts are broken in the BREAK_WRITE case to ensure that layout-holders do not collide with local writes. Additionally, layouts are broken in the BREAK_UNMAP case to make sure the layout-holder has a consistent view of the file's extent map. While BREAK_WRITE breaks can be satisfied be recalling FL_LAYOUT leases, BREAK_UNMAP breaks additionally require waiting for busy dax-pages to go idle while holding XFS_MMAPLOCK_EXCL. After this refactoring xfs_break_layouts() becomes the entry point for coordinating both types of breaks. Finally, xfs_break_leased_layouts() becomes just the BREAK_WRITE handler. Note that the unlock tracking is needed in a follow on change. That will coordinate retrying either break handler until both successfully test for a lease break while maintaining the lock state. Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Cc: "Darrick J. Wong" <darrick.wong@oracle.com> Reported-by: Dave Chinner <david@fromorbit.com> Reported-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-03-21 05:42:38 +08:00
}
/* Does this file, inode, or mount want synchronous writes? */
static inline bool xfs_file_sync_writes(struct file *filp)
{
struct xfs_inode *ip = XFS_I(file_inode(filp));
if (xfs_has_wsync(ip->i_mount))
return true;
if (filp->f_flags & (__O_SYNC | O_DSYNC))
return true;
if (IS_SYNC(file_inode(filp)))
return true;
return false;
}
#define XFS_FALLOC_FL_SUPPORTED \
(FALLOC_FL_KEEP_SIZE | FALLOC_FL_PUNCH_HOLE | \
FALLOC_FL_COLLAPSE_RANGE | FALLOC_FL_ZERO_RANGE | \
FALLOC_FL_INSERT_RANGE | FALLOC_FL_UNSHARE_RANGE)
STATIC long
xfs_file_fallocate(
struct file *file,
int mode,
loff_t offset,
loff_t len)
{
struct inode *inode = file_inode(file);
struct xfs_inode *ip = XFS_I(inode);
long error;
uint iolock = XFS_IOLOCK_EXCL | XFS_MMAPLOCK_EXCL;
loff_t new_size = 0;
bool do_file_insert = false;
if (!S_ISREG(inode->i_mode))
return -EINVAL;
if (mode & ~XFS_FALLOC_FL_SUPPORTED)
return -EOPNOTSUPP;
xfs_ilock(ip, iolock);
xfs: prepare xfs_break_layouts() for another layout type When xfs is operating as the back-end of a pNFS block server, it prevents collisions between local and remote operations by requiring a lease to be held for remotely accessed blocks. Local filesystem operations break those leases before writing or mutating the extent map of the file. A similar mechanism is needed to prevent operations on pinned dax mappings, like device-DMA, from colliding with extent unmap operations. BREAK_WRITE and BREAK_UNMAP are introduced as two distinct levels of layout breaking. Layouts are broken in the BREAK_WRITE case to ensure that layout-holders do not collide with local writes. Additionally, layouts are broken in the BREAK_UNMAP case to make sure the layout-holder has a consistent view of the file's extent map. While BREAK_WRITE breaks can be satisfied be recalling FL_LAYOUT leases, BREAK_UNMAP breaks additionally require waiting for busy dax-pages to go idle while holding XFS_MMAPLOCK_EXCL. After this refactoring xfs_break_layouts() becomes the entry point for coordinating both types of breaks. Finally, xfs_break_leased_layouts() becomes just the BREAK_WRITE handler. Note that the unlock tracking is needed in a follow on change. That will coordinate retrying either break handler until both successfully test for a lease break while maintaining the lock state. Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Cc: "Darrick J. Wong" <darrick.wong@oracle.com> Reported-by: Dave Chinner <david@fromorbit.com> Reported-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2018-03-21 05:42:38 +08:00
error = xfs_break_layouts(inode, &iolock, BREAK_UNMAP);
if (error)
goto out_unlock;
xfs: properly serialise fallocate against AIO+DIO AIO+DIO can extend the file size on IO completion, and it holds no inode locks while the IO is in flight. Therefore, a race condition exists in file size updates if we do something like this: aio-thread fallocate-thread lock inode submit IO beyond inode->i_size unlock inode ..... lock inode break layouts if (off + len > inode->i_size) new_size = off + len ..... inode_dio_wait() <blocks> ..... completes inode->i_size updated inode_dio_done() .... <wakes> <does stuff no long beyond EOF> if (new_size) xfs_vn_setattr(inode, new_size) Yup, that attempt to extend the file size in the fallocate code turns into a truncate - it removes the whatever the aio write allocated and put to disk, and reduced the inode size back down to where the fallocate operation ends. Fundamentally, xfs_file_fallocate() not compatible with racing AIO+DIO completions, so we need to move the inode_dio_wait() call up to where the lock the inode and break the layouts. Secondly, storing the inode size and then using it unchecked without holding the ILOCK is not safe; we can only do such a thing if we've locked out and drained all IO and other modification operations, which we don't do initially in xfs_file_fallocate. It should be noted that some of the fallocate operations are compound operations - they are made up of multiple manipulations that may zero data, and so we may need to flush and invalidate the file multiple times during an operation. However, we only need to lock out IO and other space manipulation operations once, as that lockout is maintained until the entire fallocate operation has been completed. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
2019-10-30 04:04:32 +08:00
/*
* Must wait for all AIO to complete before we continue as AIO can
* change the file size on completion without holding any locks we
* currently hold. We must do this first because AIO can update both
* the on disk and in memory inode sizes, and the operations that follow
* require the in-memory size to be fully up-to-date.
*/
inode_dio_wait(inode);
/*
* Now AIO and DIO has drained we flush and (if necessary) invalidate
* the cached range over the first operation we are about to run.
*
* We care about zero and collapse here because they both run a hole
* punch over the range first. Because that can zero data, and the range
* of invalidation for the shift operations is much larger, we still do
* the required flush for collapse in xfs_prepare_shift().
*
* Insert has the same range requirements as collapse, and we extend the
* file first which can zero data. Hence insert has the same
* flush/invalidate requirements as collapse and so they are both
* handled at the right time by xfs_prepare_shift().
*/
if (mode & (FALLOC_FL_PUNCH_HOLE | FALLOC_FL_ZERO_RANGE |
FALLOC_FL_COLLAPSE_RANGE)) {
error = xfs_flush_unmap_range(ip, offset, len);
if (error)
goto out_unlock;
}
error = file_modified(file);
if (error)
goto out_unlock;
if (mode & FALLOC_FL_PUNCH_HOLE) {
error = xfs_free_file_space(ip, offset, len);
if (error)
goto out_unlock;
} else if (mode & FALLOC_FL_COLLAPSE_RANGE) {
if (!xfs_is_falloc_aligned(ip, offset, len)) {
error = -EINVAL;
goto out_unlock;
}
/*
* There is no need to overlap collapse range with EOF,
* in which case it is effectively a truncate operation
*/
if (offset + len >= i_size_read(inode)) {
error = -EINVAL;
goto out_unlock;
}
new_size = i_size_read(inode) - len;
error = xfs_collapse_file_space(ip, offset, len);
if (error)
goto out_unlock;
} else if (mode & FALLOC_FL_INSERT_RANGE) {
loff_t isize = i_size_read(inode);
if (!xfs_is_falloc_aligned(ip, offset, len)) {
error = -EINVAL;
goto out_unlock;
}
/*
* New inode size must not exceed ->s_maxbytes, accounting for
* possible signed overflow.
*/
if (inode->i_sb->s_maxbytes - isize < len) {
error = -EFBIG;
goto out_unlock;
}
new_size = isize + len;
/* Offset should be less than i_size */
if (offset >= isize) {
error = -EINVAL;
goto out_unlock;
}
do_file_insert = true;
} else {
if (!(mode & FALLOC_FL_KEEP_SIZE) &&
offset + len > i_size_read(inode)) {
new_size = offset + len;
error = inode_newsize_ok(inode, new_size);
if (error)
goto out_unlock;
}
xfs: introduce an always_cow mode Add a mode where XFS never overwrites existing blocks in place. This is to aid debugging our COW code, and also put infatructure in place for things like possible future support for zoned block devices, which can't support overwrites. This mode is enabled globally by doing a: echo 1 > /sys/fs/xfs/debug/always_cow Note that the parameter is global to allow running all tests in xfstests easily in this mode, which would not easily be possible with a per-fs sysfs file. In always_cow mode persistent preallocations are disabled, and fallocate will fail when called with a 0 mode (with our without FALLOC_FL_KEEP_SIZE), and not create unwritten extent for zeroed space when called with FALLOC_FL_ZERO_RANGE or FALLOC_FL_UNSHARE_RANGE. There are a few interesting xfstests failures when run in always_cow mode: - generic/392 fails because the bytes used in the file used to test hole punch recovery are less after the log replay. This is because the blocks written and then punched out are only freed with a delay due to the logging mechanism. - xfs/170 will fail as the already fragile file streams mechanism doesn't seem to interact well with the COW allocator - xfs/180 xfs/182 xfs/192 xfs/198 xfs/204 and xfs/208 will claim the file system is badly fragmented, but there is not much we can do to avoid that when always writing out of place - xfs/205 fails because overwriting a file in always_cow mode will require new space allocation and the assumption in the test thus don't work anymore. - xfs/326 fails to modify the file at all in always_cow mode after injecting the refcount error, leading to an unexpected md5sum after the remount, but that again is expected Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
2019-02-19 01:38:49 +08:00
if (mode & FALLOC_FL_ZERO_RANGE) {
/*
* Punch a hole and prealloc the range. We use a hole
* punch rather than unwritten extent conversion for two
* reasons:
*
* 1.) Hole punch handles partial block zeroing for us.
* 2.) If prealloc returns ENOSPC, the file range is
* still zero-valued by virtue of the hole punch.
*/
unsigned int blksize = i_blocksize(inode);
trace_xfs_zero_file_space(ip);
error = xfs_free_file_space(ip, offset, len);
if (error)
goto out_unlock;
len = round_up(offset + len, blksize) -
round_down(offset, blksize);
offset = round_down(offset, blksize);
xfs: introduce an always_cow mode Add a mode where XFS never overwrites existing blocks in place. This is to aid debugging our COW code, and also put infatructure in place for things like possible future support for zoned block devices, which can't support overwrites. This mode is enabled globally by doing a: echo 1 > /sys/fs/xfs/debug/always_cow Note that the parameter is global to allow running all tests in xfstests easily in this mode, which would not easily be possible with a per-fs sysfs file. In always_cow mode persistent preallocations are disabled, and fallocate will fail when called with a 0 mode (with our without FALLOC_FL_KEEP_SIZE), and not create unwritten extent for zeroed space when called with FALLOC_FL_ZERO_RANGE or FALLOC_FL_UNSHARE_RANGE. There are a few interesting xfstests failures when run in always_cow mode: - generic/392 fails because the bytes used in the file used to test hole punch recovery are less after the log replay. This is because the blocks written and then punched out are only freed with a delay due to the logging mechanism. - xfs/170 will fail as the already fragile file streams mechanism doesn't seem to interact well with the COW allocator - xfs/180 xfs/182 xfs/192 xfs/198 xfs/204 and xfs/208 will claim the file system is badly fragmented, but there is not much we can do to avoid that when always writing out of place - xfs/205 fails because overwriting a file in always_cow mode will require new space allocation and the assumption in the test thus don't work anymore. - xfs/326 fails to modify the file at all in always_cow mode after injecting the refcount error, leading to an unexpected md5sum after the remount, but that again is expected Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
2019-02-19 01:38:49 +08:00
} else if (mode & FALLOC_FL_UNSHARE_RANGE) {
error = xfs_reflink_unshare(ip, offset, len);
if (error)
goto out_unlock;
} else {
/*
* If always_cow mode we can't use preallocations and
* thus should not create them.
*/
if (xfs_is_always_cow_inode(ip)) {
error = -EOPNOTSUPP;
goto out_unlock;
}
}
xfs: introduce an always_cow mode Add a mode where XFS never overwrites existing blocks in place. This is to aid debugging our COW code, and also put infatructure in place for things like possible future support for zoned block devices, which can't support overwrites. This mode is enabled globally by doing a: echo 1 > /sys/fs/xfs/debug/always_cow Note that the parameter is global to allow running all tests in xfstests easily in this mode, which would not easily be possible with a per-fs sysfs file. In always_cow mode persistent preallocations are disabled, and fallocate will fail when called with a 0 mode (with our without FALLOC_FL_KEEP_SIZE), and not create unwritten extent for zeroed space when called with FALLOC_FL_ZERO_RANGE or FALLOC_FL_UNSHARE_RANGE. There are a few interesting xfstests failures when run in always_cow mode: - generic/392 fails because the bytes used in the file used to test hole punch recovery are less after the log replay. This is because the blocks written and then punched out are only freed with a delay due to the logging mechanism. - xfs/170 will fail as the already fragile file streams mechanism doesn't seem to interact well with the COW allocator - xfs/180 xfs/182 xfs/192 xfs/198 xfs/204 and xfs/208 will claim the file system is badly fragmented, but there is not much we can do to avoid that when always writing out of place - xfs/205 fails because overwriting a file in always_cow mode will require new space allocation and the assumption in the test thus don't work anymore. - xfs/326 fails to modify the file at all in always_cow mode after injecting the refcount error, leading to an unexpected md5sum after the remount, but that again is expected Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
2019-02-19 01:38:49 +08:00
if (!xfs_is_always_cow_inode(ip)) {
error = xfs_alloc_file_space(ip, offset, len);
if (error)
goto out_unlock;
}
}
/* Change file size if needed */
if (new_size) {
struct iattr iattr;
iattr.ia_valid = ATTR_SIZE;
iattr.ia_size = new_size;
error = xfs_vn_setattr_size(file_mnt_user_ns(file),
file_dentry(file), &iattr);
if (error)
goto out_unlock;
}
/*
* Perform hole insertion now that the file size has been
* updated so that if we crash during the operation we don't
* leave shifted extents past EOF and hence losing access to
* the data that is contained within them.
*/
if (do_file_insert) {
error = xfs_insert_file_space(ip, offset, len);
if (error)
goto out_unlock;
}
if (xfs_file_sync_writes(file))
error = xfs_log_force_inode(ip);
out_unlock:
xfs_iunlock(ip, iolock);
return error;
}
STATIC int
xfs_file_fadvise(
struct file *file,
loff_t start,
loff_t end,
int advice)
{
struct xfs_inode *ip = XFS_I(file_inode(file));
int ret;
int lockflags = 0;
/*
* Operations creating pages in page cache need protection from hole
* punching and similar ops
*/
if (advice == POSIX_FADV_WILLNEED) {
lockflags = XFS_IOLOCK_SHARED;
xfs_ilock(ip, lockflags);
}
ret = generic_fadvise(file, start, end, advice);
if (lockflags)
xfs_iunlock(ip, lockflags);
return ret;
}
STATIC loff_t
xfs_file_remap_range(
struct file *file_in,
loff_t pos_in,
struct file *file_out,
loff_t pos_out,
loff_t len,
unsigned int remap_flags)
{
struct inode *inode_in = file_inode(file_in);
struct xfs_inode *src = XFS_I(inode_in);
struct inode *inode_out = file_inode(file_out);
struct xfs_inode *dest = XFS_I(inode_out);
struct xfs_mount *mp = src->i_mount;
loff_t remapped = 0;
xfs_extlen_t cowextsize;
int ret;
if (remap_flags & ~(REMAP_FILE_DEDUP | REMAP_FILE_ADVISORY))
return -EINVAL;
if (!xfs_has_reflink(mp))
return -EOPNOTSUPP;
if (xfs_is_shutdown(mp))
return -EIO;
/* Prepare and then clone file data. */
ret = xfs_reflink_remap_prep(file_in, pos_in, file_out, pos_out,
&len, remap_flags);
if (ret || len == 0)
return ret;
trace_xfs_reflink_remap_range(src, pos_in, len, dest, pos_out);
ret = xfs_reflink_remap_blocks(src, pos_in, dest, pos_out, len,
&remapped);
if (ret)
goto out_unlock;
/*
* Carry the cowextsize hint from src to dest if we're sharing the
* entire source file to the entire destination file, the source file
* has a cowextsize hint, and the destination file does not.
*/
cowextsize = 0;
if (pos_in == 0 && len == i_size_read(inode_in) &&
(src->i_diflags2 & XFS_DIFLAG2_COWEXTSIZE) &&
pos_out == 0 && len >= i_size_read(inode_out) &&
!(dest->i_diflags2 & XFS_DIFLAG2_COWEXTSIZE))
cowextsize = src->i_cowextsize;
ret = xfs_reflink_update_dest(dest, pos_out + len, cowextsize,
remap_flags);
if (ret)
goto out_unlock;
if (xfs_file_sync_writes(file_in) || xfs_file_sync_writes(file_out))
xfs_log_force_inode(dest);
out_unlock:
xfs_iunlock2_io_mmap(src, dest);
if (ret)
trace_xfs_reflink_remap_range_error(dest, ret, _RET_IP_);
return remapped > 0 ? remapped : ret;
}
STATIC int
xfs_file_open(
struct inode *inode,
struct file *file)
{
if (xfs_is_shutdown(XFS_M(inode->i_sb)))
return -EIO;
file->f_mode |= FMODE_NOWAIT | FMODE_BUF_RASYNC | FMODE_BUF_WASYNC;
return generic_file_open(inode, file);
}
STATIC int
xfs_dir_open(
struct inode *inode,
struct file *file)
{
struct xfs_inode *ip = XFS_I(inode);
unsigned int mode;
int error;
error = xfs_file_open(inode, file);
if (error)
return error;
/*
* If there are any blocks, read-ahead block 0 as we're almost
* certain to have the next operation be a read there.
*/
mode = xfs_ilock_data_map_shared(ip);
if (ip->i_df.if_nextents > 0)
error = xfs_dir3_data_readahead(ip, 0, 0);
xfs_iunlock(ip, mode);
return error;
}
STATIC int
xfs_file_release(
struct inode *inode,
struct file *filp)
{
return xfs_release(XFS_I(inode));
}
STATIC int
xfs_file_readdir(
struct file *file,
struct dir_context *ctx)
{
struct inode *inode = file_inode(file);
xfs_inode_t *ip = XFS_I(inode);
size_t bufsize;
/*
* The Linux API doesn't pass down the total size of the buffer
* we read into down to the filesystem. With the filldir concept
* it's not needed for correct information, but the XFS dir2 leaf
* code wants an estimate of the buffer size to calculate it's
* readahead window and size the buffers used for mapping to
* physical blocks.
*
* Try to give it an estimate that's good enough, maybe at some
* point we can change the ->readdir prototype to include the
* buffer size. For now we use the current glibc buffer size.
*/
bufsize = (size_t)min_t(loff_t, XFS_READDIR_BUFSIZE, ip->i_disk_size);
return xfs_readdir(NULL, ip, ctx, bufsize);
}
STATIC loff_t
xfs_file_llseek(
struct file *file,
loff_t offset,
int whence)
{
struct inode *inode = file->f_mapping->host;
if (xfs_is_shutdown(XFS_I(inode)->i_mount))
return -EIO;
switch (whence) {
default:
return generic_file_llseek(file, offset, whence);
case SEEK_HOLE:
offset = iomap_seek_hole(inode, offset, &xfs_seek_iomap_ops);
break;
case SEEK_DATA:
offset = iomap_seek_data(inode, offset, &xfs_seek_iomap_ops);
break;
}
if (offset < 0)
return offset;
return vfs_setpos(file, offset, inode->i_sb->s_maxbytes);
}
#ifdef CONFIG_FS_DAX
static int
xfs_dax_fault(
struct vm_fault *vmf,
enum page_entry_size pe_size,
bool write_fault,
pfn_t *pfn)
{
return dax_iomap_fault(vmf, pe_size, pfn, NULL,
(write_fault && !vmf->cow_page) ?
&xfs_dax_write_iomap_ops :
&xfs_read_iomap_ops);
}
#else
static int
xfs_dax_fault(
struct vm_fault *vmf,
enum page_entry_size pe_size,
bool write_fault,
pfn_t *pfn)
{
return 0;
}
#endif
/*
* Locking for serialisation of IO during page faults. This results in a lock
* ordering of:
*
* mmap_lock (MM)
* sb_start_pagefault(vfs, freeze)
* invalidate_lock (vfs/XFS_MMAPLOCK - truncate serialisation)
* page_lock (MM)
* i_lock (XFS - extent map serialisation)
*/
static vm_fault_t
__xfs_filemap_fault(
struct vm_fault *vmf,
enum page_entry_size pe_size,
bool write_fault)
{
struct inode *inode = file_inode(vmf->vma->vm_file);
struct xfs_inode *ip = XFS_I(inode);
vm_fault_t ret;
trace_xfs_filemap_fault(ip, pe_size, write_fault);
if (write_fault) {
sb_start_pagefault(inode->i_sb);
file_update_time(vmf->vma->vm_file);
}
if (IS_DAX(inode)) {
pfn_t pfn;
xfs_ilock(XFS_I(inode), XFS_MMAPLOCK_SHARED);
ret = xfs_dax_fault(vmf, pe_size, write_fault, &pfn);
if (ret & VM_FAULT_NEEDDSYNC)
ret = dax_finish_sync_fault(vmf, pe_size, pfn);
xfs_iunlock(XFS_I(inode), XFS_MMAPLOCK_SHARED);
} else {
if (write_fault) {
xfs_ilock(XFS_I(inode), XFS_MMAPLOCK_SHARED);
ret = iomap_page_mkwrite(vmf,
&xfs_buffered_write_iomap_ops);
xfs_iunlock(XFS_I(inode), XFS_MMAPLOCK_SHARED);
} else {
ret = filemap_fault(vmf);
}
}
if (write_fault)
sb_end_pagefault(inode->i_sb);
return ret;
}
static inline bool
xfs_is_write_fault(
struct vm_fault *vmf)
{
return (vmf->flags & FAULT_FLAG_WRITE) &&
(vmf->vma->vm_flags & VM_SHARED);
}
static vm_fault_t
xfs_filemap_fault(
struct vm_fault *vmf)
{
/* DAX can shortcut the normal fault path on write faults! */
return __xfs_filemap_fault(vmf, PE_SIZE_PTE,
IS_DAX(file_inode(vmf->vma->vm_file)) &&
xfs_is_write_fault(vmf));
}
static vm_fault_t
mm,fs,dax: change ->pmd_fault to ->huge_fault Patch series "1G transparent hugepage support for device dax", v2. The following series implements support for 1G trasparent hugepage on x86 for device dax. The bulk of the code was written by Mathew Wilcox a while back supporting transparent 1G hugepage for fs DAX. I have forward ported the relevant bits to 4.10-rc. The current submission has only the necessary code to support device DAX. Comments from Dan Williams: So the motivation and intended user of this functionality mirrors the motivation and users of 1GB page support in hugetlbfs. Given expected capacities of persistent memory devices an in-memory database may want to reduce tlb pressure beyond what they can already achieve with 2MB mappings of a device-dax file. We have customer feedback to that effect as Willy mentioned in his previous version of these patches [1]. [1]: https://lkml.org/lkml/2016/1/31/52 Comments from Nilesh @ Oracle: There are applications which have a process model; and if you assume 10,000 processes attempting to mmap all the 6TB memory available on a server; we are looking at the following: processes : 10,000 memory : 6TB pte @ 4k page size: 8 bytes / 4K of memory * #processes = 6TB / 4k * 8 * 10000 = 1.5GB * 80000 = 120,000GB pmd @ 2M page size: 120,000 / 512 = ~240GB pud @ 1G page size: 240GB / 512 = ~480MB As you can see with 2M pages, this system will use up an exorbitant amount of DRAM to hold the page tables; but the 1G pages finally brings it down to a reasonable level. Memory sizes will keep increasing; so this number will keep increasing. An argument can be made to convert the applications from process model to thread model, but in the real world that may not be always practical. Hopefully this helps explain the use case where this is valuable. This patch (of 3): In preparation for adding the ability to handle PUD pages, convert vm_operations_struct.pmd_fault to vm_operations_struct.huge_fault. The vm_fault structure is extended to include a union of the different page table pointers that may be needed, and three flag bits are reserved to indicate which type of pointer is in the union. [ross.zwisler@linux.intel.com: remove unused function ext4_dax_huge_fault()] Link: http://lkml.kernel.org/r/1485813172-7284-1-git-send-email-ross.zwisler@linux.intel.com [dave.jiang@intel.com: clear PMD or PUD size flags when in fall through path] Link: http://lkml.kernel.org/r/148589842696.5820.16078080610311444794.stgit@djiang5-desk3.ch.intel.com Link: http://lkml.kernel.org/r/148545058784.17912.6353162518188733642.stgit@djiang5-desk3.ch.intel.com Signed-off-by: Matthew Wilcox <mawilcox@microsoft.com> Signed-off-by: Dave Jiang <dave.jiang@intel.com> Signed-off-by: Ross Zwisler <ross.zwisler@linux.intel.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Jan Kara <jack@suse.com> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: Nilesh Choudhury <nilesh.choudhury@oracle.com> Cc: Ingo Molnar <mingo@elte.hu> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Dave Jiang <dave.jiang@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-02-25 06:56:59 +08:00
xfs_filemap_huge_fault(
struct vm_fault *vmf,
enum page_entry_size pe_size)
{
if (!IS_DAX(file_inode(vmf->vma->vm_file)))
return VM_FAULT_FALLBACK;
/* DAX can shortcut the normal fault path on write faults! */
return __xfs_filemap_fault(vmf, pe_size,
xfs_is_write_fault(vmf));
}
static vm_fault_t
xfs_filemap_page_mkwrite(
struct vm_fault *vmf)
{
return __xfs_filemap_fault(vmf, PE_SIZE_PTE, true);
}
/*
* pfn_mkwrite was originally intended to ensure we capture time stamp updates
* on write faults. In reality, it needs to serialise against truncate and
* prepare memory for writing so handle is as standard write fault.
*/
static vm_fault_t
xfs_filemap_pfn_mkwrite(
struct vm_fault *vmf)
{
return __xfs_filemap_fault(vmf, PE_SIZE_PTE, true);
}
static vm_fault_t
xfs_filemap_map_pages(
struct vm_fault *vmf,
pgoff_t start_pgoff,
pgoff_t end_pgoff)
{
struct inode *inode = file_inode(vmf->vma->vm_file);
vm_fault_t ret;
xfs_ilock(XFS_I(inode), XFS_MMAPLOCK_SHARED);
ret = filemap_map_pages(vmf, start_pgoff, end_pgoff);
xfs_iunlock(XFS_I(inode), XFS_MMAPLOCK_SHARED);
return ret;
}
static const struct vm_operations_struct xfs_file_vm_ops = {
.fault = xfs_filemap_fault,
mm,fs,dax: change ->pmd_fault to ->huge_fault Patch series "1G transparent hugepage support for device dax", v2. The following series implements support for 1G trasparent hugepage on x86 for device dax. The bulk of the code was written by Mathew Wilcox a while back supporting transparent 1G hugepage for fs DAX. I have forward ported the relevant bits to 4.10-rc. The current submission has only the necessary code to support device DAX. Comments from Dan Williams: So the motivation and intended user of this functionality mirrors the motivation and users of 1GB page support in hugetlbfs. Given expected capacities of persistent memory devices an in-memory database may want to reduce tlb pressure beyond what they can already achieve with 2MB mappings of a device-dax file. We have customer feedback to that effect as Willy mentioned in his previous version of these patches [1]. [1]: https://lkml.org/lkml/2016/1/31/52 Comments from Nilesh @ Oracle: There are applications which have a process model; and if you assume 10,000 processes attempting to mmap all the 6TB memory available on a server; we are looking at the following: processes : 10,000 memory : 6TB pte @ 4k page size: 8 bytes / 4K of memory * #processes = 6TB / 4k * 8 * 10000 = 1.5GB * 80000 = 120,000GB pmd @ 2M page size: 120,000 / 512 = ~240GB pud @ 1G page size: 240GB / 512 = ~480MB As you can see with 2M pages, this system will use up an exorbitant amount of DRAM to hold the page tables; but the 1G pages finally brings it down to a reasonable level. Memory sizes will keep increasing; so this number will keep increasing. An argument can be made to convert the applications from process model to thread model, but in the real world that may not be always practical. Hopefully this helps explain the use case where this is valuable. This patch (of 3): In preparation for adding the ability to handle PUD pages, convert vm_operations_struct.pmd_fault to vm_operations_struct.huge_fault. The vm_fault structure is extended to include a union of the different page table pointers that may be needed, and three flag bits are reserved to indicate which type of pointer is in the union. [ross.zwisler@linux.intel.com: remove unused function ext4_dax_huge_fault()] Link: http://lkml.kernel.org/r/1485813172-7284-1-git-send-email-ross.zwisler@linux.intel.com [dave.jiang@intel.com: clear PMD or PUD size flags when in fall through path] Link: http://lkml.kernel.org/r/148589842696.5820.16078080610311444794.stgit@djiang5-desk3.ch.intel.com Link: http://lkml.kernel.org/r/148545058784.17912.6353162518188733642.stgit@djiang5-desk3.ch.intel.com Signed-off-by: Matthew Wilcox <mawilcox@microsoft.com> Signed-off-by: Dave Jiang <dave.jiang@intel.com> Signed-off-by: Ross Zwisler <ross.zwisler@linux.intel.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Jan Kara <jack@suse.com> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: Nilesh Choudhury <nilesh.choudhury@oracle.com> Cc: Ingo Molnar <mingo@elte.hu> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Dave Jiang <dave.jiang@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-02-25 06:56:59 +08:00
.huge_fault = xfs_filemap_huge_fault,
.map_pages = xfs_filemap_map_pages,
.page_mkwrite = xfs_filemap_page_mkwrite,
.pfn_mkwrite = xfs_filemap_pfn_mkwrite,
};
STATIC int
xfs_file_mmap(
struct file *file,
struct vm_area_struct *vma)
{
struct inode *inode = file_inode(file);
struct xfs_buftarg *target = xfs_inode_buftarg(XFS_I(inode));
/*
* We don't support synchronous mappings for non-DAX files and
* for DAX files if underneath dax_device is not synchronous.
*/
if (!daxdev_mapping_supported(vma, target->bt_daxdev))
return -EOPNOTSUPP;
file_accessed(file);
vma->vm_ops = &xfs_file_vm_ops;
if (IS_DAX(inode))
dax: remove VM_MIXEDMAP for fsdax and device dax This patch is reworked from an earlier patch that Dan has posted: https://patchwork.kernel.org/patch/10131727/ VM_MIXEDMAP is used by dax to direct mm paths like vm_normal_page() that the memory page it is dealing with is not typical memory from the linear map. The get_user_pages_fast() path, since it does not resolve the vma, is already using {pte,pmd}_devmap() as a stand-in for VM_MIXEDMAP, so we use that as a VM_MIXEDMAP replacement in some locations. In the cases where there is no pte to consult we fallback to using vma_is_dax() to detect the VM_MIXEDMAP special case. Now that we have explicit driver pfn_t-flag opt-in/opt-out for get_user_pages() support for DAX we can stop setting VM_MIXEDMAP. This also means we no longer need to worry about safely manipulating vm_flags in a future where we support dynamically changing the dax mode of a file. DAX should also now be supported with madvise_behavior(), vma_merge(), and copy_page_range(). This patch has been tested against ndctl unit test. It has also been tested against xfstests commit: 625515d using fake pmem created by memmap and no additional issues have been observed. Link: http://lkml.kernel.org/r/152847720311.55924.16999195879201817653.stgit@djiang5-desk3.ch.intel.com Signed-off-by: Dave Jiang <dave.jiang@intel.com> Acked-by: Dan Williams <dan.j.williams@intel.com> Cc: Jan Kara <jack@suse.cz> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2018-08-18 06:43:40 +08:00
vma->vm_flags |= VM_HUGEPAGE;
return 0;
}
const struct file_operations xfs_file_operations = {
.llseek = xfs_file_llseek,
.read_iter = xfs_file_read_iter,
.write_iter = xfs_file_write_iter,
.splice_read = generic_file_splice_read,
.splice_write = iter_file_splice_write,
.iopoll = iocb_bio_iopoll,
.unlocked_ioctl = xfs_file_ioctl,
#ifdef CONFIG_COMPAT
.compat_ioctl = xfs_file_compat_ioctl,
#endif
.mmap = xfs_file_mmap,
.mmap_supported_flags = MAP_SYNC,
.open = xfs_file_open,
.release = xfs_file_release,
.fsync = xfs_file_fsync,
.get_unmapped_area = thp_get_unmapped_area,
.fallocate = xfs_file_fallocate,
.fadvise = xfs_file_fadvise,
.remap_file_range = xfs_file_remap_range,
};
const struct file_operations xfs_dir_file_operations = {
.open = xfs_dir_open,
.read = generic_read_dir,
.iterate_shared = xfs_file_readdir,
.llseek = generic_file_llseek,
.unlocked_ioctl = xfs_file_ioctl,
#ifdef CONFIG_COMPAT
.compat_ioctl = xfs_file_compat_ioctl,
#endif
.fsync = xfs_dir_fsync,
};