IF YOU WOULD LIKE TO GET AN ACCOUNT, please write an
email to Administrator. User accounts are meant only to access repo
and report issues and/or generate pull requests.
This is a purpose-specific Git hosting for
BaseALT
projects. Thank you for your understanding!
Только зарегистрированные пользователи имеют доступ к сервису!
Для получения аккаунта, обратитесь к администратору.
Now that the generic btree code supports overlapping intervals, plug
in the rmap btree to this functionality. We will need it to find
potential left neighbors in xfs_rmap_{alloc,free} later in the patch
set.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Originally-From: Dave Chinner <dchinner@redhat.com>
Implement the generic btree operations needed to manipulate rmap
btree blocks. This is very similar to the per-ag freespace btree
implementation, and uses the AGFL for allocation and freeing of
blocks.
Adapt the rmap btree to store owner offsets within each rmap record,
and to handle the primary key being redefined as the tuple
[agblk, owner, offset]. The expansion of the primary key is crucial
to allowing multiple owners per extent.
[darrick: adapt the btree ops to deal with offsets]
[darrick: remove init_rec_from_key]
[darrick: move unwritten bit to rm_offset]
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Originally-From: Dave Chinner <dchinner@redhat.com>
The rmap btree is allocated from the AGFL, which means we have to
ensure ENOSPC is reported to userspace before we run out of free
space in each AG. The last allocation in an AG can cause a full
height rmap btree split, and that means we have to reserve at least
this many blocks *in each AG* to be placed on the AGFL at ENOSPC.
Update the various space calculation functions to handle this.
Also, because the macros are now executing conditional code and are
called quite frequently, convert them to functions that initialise
variables in the struct xfs_mount, use the new variables everywhere
and document the calculations better.
[darrick.wong@oracle.com: don't reserve blocks if !rmap]
[dchinner@redhat.com: update m_ag_max_usable after growfs]
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
The rmap btrees will use the AGFL as the block allocation source, so
we need to ensure that the transaction reservations reflect the fact
this tree is modified by allocation and freeing. Hence we need to
extend all the extent allocation/free reservations used in
transactions to handle this.
Note that this also gets rid of the unused XFS_ALLOCFREE_LOG_RES
macro, as we now do buffer reservations based on the number of
buffers logged via xfs_calc_buf_res(). Hence we only need the buffer
count calculation now.
[darrick: use rmap_maxlevels when calculating log block resv]
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Originally-From: Dave Chinner <dchinner@redhat.com>
Now we can read and write rmap btree blocks, we can add support to
the growfs code to initialise new rmap btree blocks.
[darrick.wong@oracle.com: fill out the rmap offset fields]
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Originally-From: Dave Chinner <dchinner@redhat.com>
Now we have all the surrounding call infrastructure in place, we can
start filling out the rmap btree implementation. Start with the
on-disk btree format; add everything needed to read, write and
manipulate rmap btree blocks. This prepares the way for adding the
btree operations implementation.
[darrick: record owner and offset info in rmap btree]
[darrick: fork, bmbt and unwritten state in rmap btree]
[darrick: flags are a separate field in xfs_rmap_irec]
[darrick: calculate maxlevels separately]
[darrick: move the 'unwritten' bit into unused parts of rm_offset]
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Originally-From: Dave Chinner <dchinner@redhat.com>
Add the stubs into the extent allocation and freeing paths that the
rmap btree implementation will hook into. While doing this, add the
trace points that will be used to track rmap btree extent
manipulations.
[darrick.wong@oracle.com: Extend the stubs to take full owner info.]
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
For the rmap btree to work, we have to feed the extent owner
information to the the allocation and freeing functions. This
information is what will end up in the rmap btree that tracks
allocated extents. While we technically don't need the owner
information when freeing extents, passing it allows us to validate
that the extent we are removing from the rmap btree actually
belonged to the owner we expected it to belong to.
We also define a special set of owner values for internal metadata
that would otherwise have no owner. This allows us to tell the
difference between metadata owned by different per-ag btrees, as
well as static fs metadata (e.g. AG headers) and internal journal
blocks.
There are also a couple of special cases we need to take care of -
during EFI recovery, we don't actually know who the original owner
was, so we need to pass a wildcard to indicate that we aren't
checking the owner for validity. We also need special handling in
growfs, as we "free" the space in the last AG when extending it, but
because it's new space it has no actual owner...
While touching the xfs_bmap_add_free() function, re-order the
parameters to put the struct xfs_mount first.
Extend the owner field to include both the owner type and some sort
of index within the owner. The index field will be used to support
reverse mappings when reflink is enabled.
When we're freeing extents from an EFI, we don't have the owner
information available (rmap updates have their own redo items).
xfs_free_extent therefore doesn't need to do an rmap update. Make
sure that the log replay code signals this correctly.
This is based upon a patch originally from Dave Chinner. It has been
extended to add more owner information with the intent of helping
recovery operations when things go wrong (e.g. offset of user data
block in a file).
[dchinner: de-shout the xfs_rmap_*_owner helpers]
[darrick: minor style fixes suggested by Christoph Hellwig]
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Originally-From: Dave Chinner <dchinner@redhat.com>
XFS reserves a small amount of space in each AG for the minimum
number of free blocks needed for operation. Adding the rmap btree
increases the number of reserved blocks, but it also increases the
complexity of the calculation as the free inode btree is optional
(like the rmbt).
Rather than calculate the prealloc blocks every time we need to
check it, add a function to calculate it at mount time and store it
in the struct xfs_mount, and convert the XFS_PREALLOC_BLOCKS macro
just to use the xfs-mount variable directly.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Originally-From: Dave Chinner <dchinner@redhat.com>
The rmap btree will require the same stats as all the other generic
btrees, so add all the code for that now.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Originally-From: Dave Chinner <dchinner@redhat.com>
Add new per-ag rmap btree definitions to the per-ag structures. The
rmap btree will sit in the empty slots on disk after the free space
btrees, and hence form a part of the array of space management
btrees. This requires the definition of the btree to be contiguous
with the free space btrees.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
By my calculations, a 1,073,741,824 block AG with a 1k block size
can attain a maximum height of 9. Assuming a record size of 24
bytes, a key/ptr size of 44 bytes, and half-full btree nodes, we'd
need 53,687,092 blocks for the records and ~6 million blocks for the
keys. That requires a btree of height 9 based on the following
derivation:
Block size = 1024b
sblock CRC header = 56b
== 1024-56 = 968 bytes for tree data
rmapbt record = 24b
== 40 records per leaf block
rmapbt ptr/key = 44b
== 22 ptr/keys per block
Worst case, each block is half full, so 20 records and 11 ptrs per block.
1073741824 rmap records / 20 records per block
== 53687092 leaf blocks
53687092 leaves / 11 ptrs per block
== 4880645 level 1 blocks
== 443695 level 2 blocks
== 40336 level 3 blocks
== 3667 level 4 blocks
== 334 level 5 blocks
== 31 level 6 blocks
== 3 level 7 blocks
== 1 level 8 block
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Add a couple of tracepoints for the deferred extent free operation and
a site for injecting errors while finishing the operation. This makes
it easier to debug deferred ops and test log redo.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Refactor the EFI intent item recovery (and cancellation) functions
into a general function that scans the AIL and an intent item type
specific handler. Move the function that recovers a single EFI item
into the extent free item code. We'll want the generalized function
when we start wiring up more redo item types.
Furthermore, ensure that log recovery only replays the redo items
that were in the AIL prior to recovery by checking the item LSN
against the largest LSN seen during log scanning. As written this
should never happen, but we can be defensive anyway.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Mechanical change of flist/free_list to dfops, since they're now
deferred ops, not just a freeing list.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Drop the compatibility shims that we were using to integrate the new
deferred operation mechanism into the existing code. No new code.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Restructure everything that used xfs_bmap_free to use xfs_defer_ops
instead. For now we'll just remove the old symbols and play some
cpp magic to make it work; in the next patch we'll actually rename
everything.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Connect the xfs_defer mechanism with the pieces that we'll need to
handle deferred extent freeing. We'll wire up the existing code to
our new deferred mechanism later.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Replace structure typedefs with struct xfs_foo_* in the EFI/EFD
handling code in preparation to move it over to deferred ops.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Add tracepoints for the internals of the deferred ops mechanism
and tracepoint classes for clients of the dops, to make debugging
easier.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
All the code around struct xfs_bmap_free basically implements a
deferred operation framework through which we can roll transactions
(to unlock buffers and avoid violating lock order rules) while
managing all the necessary log redo items. Previously we only used
this code to free extents after some sort of mapping operation, but
with the advent of rmap and reflink, we suddenly need to do more than
that.
With that in mind, xfs_bmap_free really becomes a deferred ops control
structure. Rename the structure and move the deferred ops into their
own file to avoid further bloating of the bmap code.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Refactor the btree_change_owner function into a more generic apparatus
which visits all blocks in a btree. We'll use this in a subsequent
patch for counting btree blocks for AG reservations.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Create a function to enable querying of btree records mapping to a
range of keys. This will be used in subsequent patches to allow
querying the reverse mapping btree to find the extents mapped to a
range of physical blocks, though the generic code can be used for
any range query.
The overlapped query range function needs to use the btree get_block
helper because the root block could be an inode, in which case
bc_bufs[nlevels-1] will be NULL.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Dave Chinner <david@fromorbit.com>
On a filesystem with both reflink and reverse mapping enabled, it's
possible to have multiple rmap records referring to the same blocks on
disk. When overlapping intervals are possible, querying a classic
btree to find all records intersecting a given interval is inefficient
because we cannot use the left side of the search interval to filter
out non-matching records the same way that we can use the existing
btree key to filter out records coming after the right side of the
search interval. This will become important once we want to use the
rmap btree to rebuild BMBTs, or implement the (future) fsmap ioctl.
(For the non-overlapping case, we can perform such queries trivially
by starting at the left side of the interval and walking the tree
until we pass the right side.)
Therefore, extend the btree code to come closer to supporting
intervals as a first-class record attribute. This involves widening
the btree node's key space to store both the lowest key reachable via
the node pointer (as the btree does now) and the highest key reachable
via the same pointer and teaching the btree modifying functions to
keep the highest-key records up to date.
This behavior can be turned on via a new btree ops flag so that btrees
that cannot store overlapping intervals don't pay the overhead costs
in terms of extra code and disk format changes.
When we're deleting a record in a btree that supports overlapped
interval records and the deletion results in two btree blocks being
joined, we defer updating the high/low keys until after all possible
joining (at higher levels in the tree) have finished. At this point,
the btree pointers at all levels have been updated to remove the empty
blocks and we can update the low and high keys.
When we're doing this, we must be careful to update the keys of all
node pointers up to the root instead of stopping at the first set of
keys that don't need updating. This is because it's possible for a
single deletion to cause joining of multiple levels of tree, and so
we need to update everything going back to the root.
The diff_two_keys functions return < 0, 0, or > 0 if key1 is less than,
equal to, or greater than key2, respectively. This is consistent
with the rest of the kernel and the C library.
In btree_updkeys(), we need to evaluate the force_all parameter before
running the key diff to avoid reading uninitialized memory when we're
forcing a key update. This happens when we've allocated an empty slot
at level N + 1 to point to a new block at level N and we're in the
process of filling out the new keys.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Add some function pointers to bc_ops to get the btree keys for
leaf and node blocks, and to update parent keys of a block.
Convert the _btree_updkey calls to use our new pointer, and
modify the tree shape changing code to call the appropriate
get_*_keys pointer instead of _btree_copy_keys because the
overlapping btree has to calculate high key values.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
When a btree block has to be split, we pass the new block's ptr from
xfs_btree_split() back to xfs_btree_insert() via a pointer parameter;
however, we pass the block's key through the cursor's record. It is a
little weird to "initialize" a record from a key since the non-key
attributes will have garbage values.
When we go to add support for interval queries, we have to be able to
pass the lowest and highest keys accessible via a pointer. There's no
clean way to pass this back through the cursor's record field.
Therefore, pass the key directly back to xfs_btree_insert() the same
way that we pass the btree_ptr.
As a bonus, we no longer need init_rec_from_key and can drop it from the
codebase.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Dave Chinner <david@fromorbit.com>
If we make the inode root block of a btree unfull by expanding the
root, we must set *stat to 1 to signal success, rather than leaving
it uninitialized.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Dave Chinner <david@fromorbit.com>
When we're deleting realtime extents, we need to lock the summary
inode in case we need to update the summary info to prevent an assert
on the rsumip inode lock on a debug kernel. While we're at it, fix
the locking annotations so that we avoid triggering lockdep warnings.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Apparently cris doesn't require structure stride to align with the
largest type in the struct, so list[0] isn't at offset 4 like it is
everywhere else. Fix this... insofar as existing XFSes on cris are
screwed.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Dave Chinner <david@fromorbit.com>
When we're iterating inode xattrs by handle, we have to copy the
cursor back to userspace so that a subsequent invocation actually
retrieves subsequent contents.
Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Been around for long enough now, hasn't caused any regression test
failures in the past 3 months, so it's time to make it a fully
supported feature.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Eric Sandeen <sandeen@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
In xfs_finish_page_writeback(), we have a loop that looks like this:
do {
if (off < bvec->bv_offset)
goto next_bh;
if (off > end)
break;
bh->b_end_io(bh, !error);
next_bh:
off += bh->b_size;
} while ((bh = bh->b_this_page) != head);
The b_end_io function is end_buffer_async_write(), which will call
end_page_writeback() once all the buffers have marked as no longer
under IO. This issue here is that the only thing currently
protecting both the bufferhead chain and the page from being
reclaimed is the PageWriteback state held on the page.
While we attempt to limit the loop to just the buffers covered by
the IO, we still read from the buffer size and follow the next
pointer in the bufferhead chain. There is no guarantee that either
of these are valid after the PageWriteback flag has been cleared.
Hence, loops like this are completely unsafe, and result in
use-after-free issues. One such problem was caught by Calvin Owens
with KASAN:
.....
INFO: Freed in 0x103fc80ec age=18446651500051355200 cpu=2165122683 pid=-1
free_buffer_head+0x41/0x90
__slab_free+0x1ed/0x340
kmem_cache_free+0x270/0x300
free_buffer_head+0x41/0x90
try_to_free_buffers+0x171/0x240
xfs_vm_releasepage+0xcb/0x3b0
try_to_release_page+0x106/0x190
shrink_page_list+0x118e/0x1a10
shrink_inactive_list+0x42c/0xdf0
shrink_zone_memcg+0xa09/0xfa0
shrink_zone+0x2c3/0xbc0
.....
Call Trace:
<IRQ> [<ffffffff81e8b8e4>] dump_stack+0x68/0x94
[<ffffffff8153a995>] print_trailer+0x115/0x1a0
[<ffffffff81541174>] object_err+0x34/0x40
[<ffffffff815436e7>] kasan_report_error+0x217/0x530
[<ffffffff81543b33>] __asan_report_load8_noabort+0x43/0x50
[<ffffffff819d651f>] xfs_destroy_ioend+0x3bf/0x4c0
[<ffffffff819d69d4>] xfs_end_bio+0x154/0x220
[<ffffffff81de0c58>] bio_endio+0x158/0x1b0
[<ffffffff81dff61b>] blk_update_request+0x18b/0xb80
[<ffffffff821baf57>] scsi_end_request+0x97/0x5a0
[<ffffffff821c5558>] scsi_io_completion+0x438/0x1690
[<ffffffff821a8d95>] scsi_finish_command+0x375/0x4e0
[<ffffffff821c3940>] scsi_softirq_done+0x280/0x340
Where the access is occuring during IO completion after the buffer
had been freed from direct memory reclaim.
Prevent use-after-free accidents in this end_io processing loop by
pre-calculating the loop conditionals before calling bh->b_end_io().
The loop is already limited to just the bufferheads covered by the
IO in progress, so the offset checks are sufficient to prevent
accessing buffers in the chain after end_page_writeback() has been
called by the the bh->b_end_io() callout.
Yet another example of why Bufferheads Must Die.
cc: <stable@vger.kernel.org> # 4.7
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reported-and-Tested-by: Calvin Owens <calvinowens@fb.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
One of the problems we currently have with delayed logging is that
under serious memory pressure we can deadlock memory reclaim. THis
occurs when memory reclaim (such as run by kswapd) is reclaiming XFS
inodes and issues a log force to unpin inodes that are dirty in the
CIL.
The CIL is pushed, but this will only occur once it gets the CIL
context lock to ensure that all committing transactions are complete
and no new transactions start being committed to the CIL while the
push switches to a new context.
The deadlock occurs when the CIL context lock is held by a
committing process that is doing memory allocation for log vector
buffers, and that allocation is then blocked on memory reclaim
making progress. Memory reclaim, however, is blocked waiting for
a log force to make progress, and so we effectively deadlock at this
point.
To solve this problem, we have to move the CIL log vector buffer
allocation outside of the context lock so that memory reclaim can
always make progress when it needs to force the log. The problem
with doing this is that a CIL push can take place while we are
determining if we need to allocate a new log vector buffer for
an item and hence the current log vector may go away without
warning. That means we canot rely on the existing log vector being
present when we finally grab the context lock and so we must have a
replacement buffer ready to go at all times.
To ensure this, introduce a "shadow log vector" buffer that is
always guaranteed to be present when we gain the CIL context lock
and format the item. This shadow buffer may or may not be used
during the formatting, but if the log item does not have an existing
log vector buffer or that buffer is too small for the new
modifications, we swap it for the new shadow buffer and format
the modifications into that new log vector buffer.
The result of this is that for any object we modify more than once
in a given CIL checkpoint, we double the memory required
to track dirty regions in the log. For single modifications then
we consume the shadow log vectorwe allocate on commit, and that gets
consumed by the checkpoint. However, if we make multiple
modifications, then the second transaction commit will allocate a
shadow log vector and hence we will end up with double the memory
usage as only one of the log vectors is consumed by the CIL
checkpoint. The remaining shadow vector will be freed when th elog
item is freed.
This can probably be optimised in future - access to the shadow log
vector is serialised by the object lock (as opposited to the active
log vector, which is controlled by the CIL context lock) and so we
can probably free shadow log vector from some objects when the log
item is marked clean on removal from the AIL.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
xfsprogs source commit 4280e59dcbc4cd8e01585efe788a68eb378048e8
xfs_da3_split() has to handle all three versions of the
directory/attribute btree structure. The attr tree is v1, the dir
tre is v2 or v3. The main difference between the v1 and v2/3 trees
is the way tree nodes are split - in the v1 tree we can require a
double split to occur because the object to be inserted may be
larger than the space made by splitting a leaf. In this case we need
to do a double split - one to split the full leaf, then another to
allocate an empty leaf block in the correct location for the new
entry. This does not happen with dir (v2/v3) formats as the objects
being inserted are always guaranteed to fit into the new space in
the split blocks.
Indeed, for directories they *may* be an extra block on this buffer
pointer. However, it's guaranteed not to be a leaf block (i.e. a
directory data block) - the directory code only ever places hash
index or free space blocks in this pointer (as a cursor of
sorts), and so to use it as a directory data block will immediately
corrupt the directory.
The problem is that the code assumes that there may be extra blocks
that we need to link into the tree once we've split the root, but
this is not true for either dir or attr trees, because the extra
attr block is always consumed by the last node split before we split
the root. Hence the linking in an extra block is always wrong at the
root split level, and this manifests itself in repair as a directory
corruption in a repaired directory, leaving the directory rebuild
incomplete.
This is a dir v2 zero-day bug - it was in the initial dir v2 commit
that was made back in February 1998.
Fix this by ensuring the linking of the blocks after the root split
never tries to make use of the extra blocks that may be held in the
cursor. They are held there for other purposes and should never be
touched by the root splitting code.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Brian Foster <bfoster@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Dave Chinner <david@fromorbit.com>
We check IS_DAX(inode) before calling either xfs_file_dax_read or
xfs_file_dax_write, and this will lead the call being optimized out at
compile time when CONFIG_FS_DAX is disabled.
However, the two functions are marked STATIC, so they become global
symbols when CONFIG_XFS_DEBUG is set, leaving us with two unused global
functions that call into an undefined function and a broken "allmodconfig"
build:
fs/built-in.o: In function `xfs_file_dax_read':
fs/xfs/xfs_file.c:348: undefined reference to `dax_do_io'
fs/built-in.o: In function `xfs_file_dax_write':
fs/xfs/xfs_file.c:758: undefined reference to `dax_do_io'
Marking the two functions 'static noinline' instead of 'STATIC' will let
the compiler drop the symbols when there are no callers but avoid the
implicit inlining.
Signed-off-by: Arnd Bergmann <arnd@arndb.de>
Fixes: 16d4d43595 ("xfs: split direct I/O and DAX path")
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Dave Chinner <david@fromorbit.com>
XFS has had scattered reports of delalloc blocks present at
->releasepage() time. This results in a warning with a stack trace
similar to the following:
...
Call Trace:
[<ffffffffa23c5b8f>] dump_stack+0x63/0x84
[<ffffffffa20837a7>] warn_slowpath_common+0x97/0xe0
[<ffffffffa208380a>] warn_slowpath_null+0x1a/0x20
[<ffffffffa2326caf>] xfs_vm_releasepage+0x10f/0x140
[<ffffffffa218c680>] ? page_mkclean_one+0xd0/0xd0
[<ffffffffa218d3a0>] ? anon_vma_prepare+0x150/0x150
[<ffffffffa21521c2>] try_to_release_page+0x32/0x50
[<ffffffffa2166b2e>] shrink_active_list+0x3ce/0x3e0
[<ffffffffa21671c7>] shrink_lruvec+0x687/0x7d0
[<ffffffffa21673ec>] shrink_zone+0xdc/0x2c0
[<ffffffffa2168539>] kswapd+0x4f9/0x970
[<ffffffffa2168040>] ? mem_cgroup_shrink_node_zone+0x1a0/0x1a0
[<ffffffffa20a0d99>] kthread+0xc9/0xe0
[<ffffffffa20a0cd0>] ? kthread_stop+0x100/0x100
[<ffffffffa26b404f>] ret_from_fork+0x3f/0x70
[<ffffffffa20a0cd0>] ? kthread_stop+0x100/0x100
This occurs because it is possible for shrink_active_list() to send
pages marked dirty to ->releasepage() when certain buffer_head threshold
conditions are met. shrink_active_list() doesn't check the page dirty
state apparently to handle an old ext3 corner case where in some cases
clean pages would not have the dirty bit cleared, thus it is up to the
filesystem to determine how to handle the page.
XFS currently handles the delalloc case properly, but this behavior
makes the warning spurious. Update the XFS ->releasepage() handler to
explicitly skip dirty pages. Retain the existing delalloc/unwritten
checks so we continue to warn if such buffers exist on clean pages when
they shouldn't.
Diagnosed-by: Dave Chinner <david@fromorbit.com>
Signed-off-by: Brian Foster <bfoster@redhat.com>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Instead we always declare struct xfs_dir2_sf_hdr as packed. That's
the expected layout, and while most major architectures do the packing
by default the new structure size and offset checker showed that not
only the ARM old ABI got this wrong, but various minor embedded
architectures did as well.
[Verified that no code change on x86-64 results from this change]
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
And use an array of unsigned char values directly to avoid problems
with architectures that pad the size of structures. This also gets
rid of the xfs_dir2_ino4_t and xfs_dir2_ino8_t types, and introduces
new constants for the size of 4 and 8 bytes as well as the size
difference between the two.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Just use an array of two unsigned chars directly to avoid problems
with architectures that pad the size of structures.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
So far the DAX code overloaded the direct I/O code path. There is very little
in common between the two, and untangling them allows to clean up both variants.
As a side effect we also get separate trace points for both I/O types.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
We control both the callers and callees of ->direct_IO, so remove the
indirect calls.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
XFS already implement it's own flushing of the pagecache because it
implements proper synchronization for direct I/O reads. This means
calling generic_file_read_iter for direct I/O is rather useless,
as it doesn't do much but updating the atime and iocb position for
us. This also gets rid of the buffered I/O fallback that isn't used
for XFS.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Similar to what we did on the write side a while ago.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
All the three low-level read implementations that we might call already
take care of not overflowing the maximum supported bytes, no need to
duplicate it here.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
Now that we have the direct I/O kiocb flag there is no real need to sample
the value inside of XFS, and the invis flag was always just partially used
and isn't worth keeping this infrastructure around for. This also splits
the read tracepoint into buffered vs direct as we've done for writes a long
time ago.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Dave Chinner <david@fromorbit.com>
