- 20 Sep, 2017 40 commits
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Greg Kroah-Hartman authored
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Darrick J. Wong authored
commit 7bf7a193 upstream. Fix up all the compiler warnings that have crept in. Signed-off-by:
Darrick J. Wong <darrick.wong@oracle.com> Reviewed-by:
Christoph Hellwig <hch@lst.de> Cc: Arnd Bergmann <arnd@arndb.de> Signed-off-by:
Greg Kroah-Hartman <gregkh@linuxfoundation.org>
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Song Liu authored
commit 9c72a18e upstream. In raid5, there are scenarios where some ios are deferred to a later time, and some IO need a flush to complete. To make sure we make progress with these IOs, we need to call the following functions: flush_deferred_bios(conf); r5l_flush_stripe_to_raid(conf->log); Both of these functions are called in raid5d(), but missing in raid5_do_work(). As a result, these functions are not called when multi-threading (group_thread_cnt > 0) is enabled. This patch adds calls to these function to raid5_do_work(). Note for stable branches: r5l_flush_stripe_to_raid(conf->log) is need for 4.4+ flush_deferred_bios(conf) is only needed for 4.11+ Signed-off-by:
Song Liu <songliubraving@fb.com> Signed-off-by:
Shaohua Li <shli@fb.com> Signed-off-by:
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Shaohua Li authored
commit 208410b5 upstream. Data allocated from mempool doesn't always get initialized, this happens when the data is reused instead of fresh allocation. In the raid1/10 case, we must reinitialize the bios. Reported-by:
Jonathan G. Underwood <jonathan.underwood@gmail.com> Fixes: f0250618(md: raid10: don't use bio's vec table to manage resync pages) Fixes: 98d30c58(md: raid1: don't use bio's vec table to manage resync pages) Cc: Ming Lei <ming.lei@redhat.com> Signed-off-by:
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Pan Bian authored
commit 6c370590 upstream. In function xfs_test_remount_options(), kfree() is used to free memory allocated by kmem_zalloc(). But it is better to use kmem_free(). Signed-off-by:
Pan Bian <bianpan2016@163.com> Reviewed-by:
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Christoph Hellwig authored
commit 8353a814 upstream. Our loop in xfs_finish_page_writeback, which iterates over all buffer heads in a page and then calls end_buffer_async_write, which also iterates over all buffers in the page to check if any I/O is in flight is not only inefficient, but also potentially dangerous as end_buffer_async_write can cause the page and all buffers to be freed. Replace it with a single loop that does the work of end_buffer_async_write on a per-page basis. Signed-off-by:
Brian Foster <bfoster@redhat.com> Reviewed-by:
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Christoph Hellwig authored
commit dd60687e upstream. Reject attempts to set XFLAGS that correspond to di_flags2 inode flags if the inode isn't a v3 inode, because di_flags2 only exists on v3. Signed-off-by:
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Amir Goldstein authored
commit 47c7d0b1 upstream. When calling into _xfs_log_force{,_lsn}() with a pointer to log_flushed variable, log_flushed will be set to 1 if: 1. xlog_sync() is called to flush the active log buffer AND/OR 2. xlog_wait() is called to wait on a syncing log buffers xfs_file_fsync() checks the value of log_flushed after _xfs_log_force_lsn() call to optimize away an explicit PREFLUSH request to the data block device after writing out all the file's pages to disk. This optimization is incorrect in the following sequence of events: Task A Task B ------------------------------------------------------- xfs_file_fsync() _xfs_log_force_lsn() xlog_sync() [submit PREFLUSH] xfs_file_fsync() file_write_and_wait_range() [submit WRITE X] [endio WRITE X] _xfs_log_force_lsn() xlog_wait() [endio PREFLUSH] The write X is not guarantied to be on persistent storage when PREFLUSH request in completed, because write A was submitted after the PREFLUSH request, but xfs_file_fsync() of task A will be notified of log_flushed=1 and will skip explicit flush. If the system crashes after fsync of task A, write X may not be present on disk after reboot. This bug was discovered and demonstrated using Josef Bacik's dm-log-writes target, which can be used to record block io operations and then replay a subset of these operations onto the target device. The test goes something like this: - Use fsx to execute ops of a file and record ops on log device - Every now and then fsync the file, store md5 of file and mark the location in the log - Then replay log onto device for each mark, mount fs and compare md5 of file to stored value Cc: Christoph Hellwig <hch@lst.de> Cc: Josef Bacik <jbacik@fb.com> Cc: <stable@vger.kernel.org> Signed-off-by:
Amir Goldstein <amir73il@gmail.com> Reviewed-by:
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Christoph Hellwig authored
commit 742d8429 upstream. Currently flag switching can be used to easily crash the kernel. Disable the per-inode DAX flag until that is sorted out. Signed-off-by:
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Brian Foster authored
commit 2dd3d709 upstream. The owner change bmbt scan that occurs during extent swap operations does not handle ordered buffer failures. Buffers that cannot be marked ordered must be physically logged so previously dirty ranges of the buffer can be relogged in the transaction. Since the bmbt scan may need to process and potentially log a large number of blocks, we can't expect to complete this operation in a single transaction. Update extent swap to use a permanent transaction with enough log reservation to physically log a buffer. Update the bmbt scan to physically log any buffers that cannot be ordered and to terminate the scan with -EAGAIN. On -EAGAIN, the caller rolls the transaction and restarts the scan. Finally, update the bmbt scan helper function to skip bmbt blocks that already match the expected owner so they are not reprocessed after scan restarts. Signed-off-by:
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Brian Foster authored
commit a5814bce upstream. Ordered buffers are used in situations where the buffer is not physically logged but must pass through the transaction/logging pipeline for a particular transaction. As a result, ordered buffers are not unpinned and written back until the transaction commits to the log. Ordered buffers have a strict requirement that the target buffer must not be currently dirty and resident in the log pipeline at the time it is marked ordered. If a dirty+ordered buffer is committed, the buffer is reinserted to the AIL but not physically relogged at the LSN of the associated checkpoint. The buffer log item is assigned the LSN of the latest checkpoint and the AIL effectively releases the previously logged buffer content from the active log before the buffer has been written back. If the tail pushes forward and a filesystem crash occurs while in this state, an inconsistent filesystem could result. It is currently the caller responsibility to ensure an ordered buffer is not already dirty from a previous modification. This is unclear and error prone when not used in situations where it is guaranteed a buffer has not been previously modified (such as new metadata allocations). To facilitate general purpose use of ordered buffers, update xfs_trans_ordered_buf() to conditionally order the buffer based on state of the log item and return the status of the result. If the bli is dirty, do not order the buffer and return false. The caller must either physically log the buffer (having acquired the appropriate log reservation) or push it from the AIL to clean it before it can be marked ordered in the current transaction. Note that ordered buffers are currently only used in two situations: 1.) inode chunk allocation where previously logged buffers are not possible and 2.) extent swap which will be updated to handle ordered buffer failures in a separate patch. Signed-off-by:
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Brian Foster authored
commit 6fb10d6d upstream. The extent swap operation currently resets bmbt block owners before the inode forks are swapped. The bmbt buffers are marked as ordered so they do not have to be physically logged in the transaction. This use of ordered buffers is not safe as bmbt buffers may have been previously physically logged. The bmbt owner change algorithm needs to be updated to physically log buffers that are already dirty when/if they are encountered. This means that an extent swap will eventually require multiple rolling transactions to handle large btrees. In addition, all inode related changes must be logged before the bmbt owner change scan begins and can roll the transaction for the first time to preserve fs consistency via log recovery. In preparation for such fixes to the bmbt owner change algorithm, refactor the bmbt scan out of the extent fork swap code to the last operation before the transaction is committed. Update xfs_swap_extent_forks() to only set the inode log flags when an owner change scan is necessary. Update xfs_swap_extents() to trigger the owner change based on the inode log flags. Note that since the owner change now occurs after the extent fork swap, the inode btrees must be fixed up with the inode number of the current inode (similar to log recovery). Signed-off-by:
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Brian Foster authored
commit 99c794c6 upstream. Extent swap uses xfs_btree_visit_blocks() to fix up bmbt block owners on v5 (!rmapbt) filesystems. The bmbt scan uses xfs_btree_lookup_get_block() to read bmbt blocks which verifies the current owner of the block against the parent inode of the bmbt. This works during extent swap because the bmbt owners are updated to the opposite inode number before the inode extent forks are swapped. The modified bmbt blocks are marked as ordered buffers which allows everything to commit in a single transaction. If the transaction commits to the log and the system crashes such that recovery of the extent swap is required, log recovery restarts the bmbt scan to fix up any bmbt blocks that may have not been written back before the crash. The log recovery bmbt scan occurs after the inode forks have been swapped, however. This causes the bmbt block owner verification to fail, leads to log recovery failure and requires xfs_repair to zap the log to recover. Define a new invalid inode owner flag to inform the btree block lookup mechanism that the current inode may be invalid with respect to the current owner of the bmbt block. Set this flag on the cursor used for change owner scans to allow this operation to work at runtime and during log recovery. Signed-off-by:
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Brian Foster authored
commit 8dc518df upstream. Ordered buffers are attached to transactions and pushed through the logging infrastructure just like normal buffers with the exception that they are not actually written to the log. Therefore, we don't need to log dirty ranges of ordered buffers. xfs_trans_log_buf() is called on ordered buffers to set up all of the dirty state on the transaction, buffer and log item and prepare the buffer for I/O. Now that xfs_trans_dirty_buf() is available, call it from xfs_trans_ordered_buf() so the latter is now mutually exclusive with xfs_trans_log_buf(). This reflects the implementation of ordered buffers and helps eliminate confusion over the need to log ranges of ordered buffers just to set up internal log state. Signed-off-by:
Allison Henderson <allison.henderson@oracle.com> Reviewed-by:
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Brian Foster authored
commit 9684010d upstream. xfs_trans_log_buf() is responsible for logging the dirty segments of a buffer along with setting all of the necessary state on the transaction, buffer, bli, etc., to ensure that the associated items are marked as dirty and prepared for I/O. We have a couple use cases that need to to dirty a buffer in a transaction without actually logging dirty ranges of the buffer. One existing use case is ordered buffers, which are currently logged with arbitrary ranges to accomplish this even though the content of ordered buffers is never written to the log. Another pending use case is to relog an already dirty buffer across rolled transactions within the deferred operations infrastructure. This is required to prevent a held (XFS_BLI_HOLD) buffer from pinning the tail of the log. Refactor xfs_trans_log_buf() into a new function that contains all of the logic responsible to dirty the transaction, lidp, buffer and bli. This new function can be used in the future for the use cases outlined above. This patch does not introduce functional changes. Signed-off-by:
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Brian Foster authored
commit e9385cc6 upstream. Ordered buffers pass through the logging infrastructure without ever being written to the log. The way this works is that the ordered buffer status is transferred to the log vector at commit time via the ->iop_size() callback. In xlog_cil_insert_format_items(), ordered log vectors bypass ->iop_format() processing altogether. Therefore it is unnecessary for xfs_buf_item_format() to handle ordered buffers. Remove the unnecessary logic and assert that an ordered buffer never reaches this point. Signed-off-by:
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Brian Foster authored
commit 6453c65d upstream. xfs_buf_item_unlock() historically checked the dirty state of the buffer by manually checking the buffer log formats for dirty segments. The introduction of ordered buffers invalidated this check because ordered buffers have dirty bli's but no dirty (logged) segments. The check was updated to accommodate ordered buffers by looking at the bli state first and considering the blf only if the bli is clean. This logic is safe but unnecessary. There is no valid case where the bli is clean yet the blf has dirty segments. The bli is set dirty whenever the blf is logged (via xfs_trans_log_buf()) and the blf is cleared in the only place BLI_DIRTY is cleared (xfs_trans_binval()). Remove the conditional blf dirty checks and replace with an assert that should catch any discrepencies between bli and blf dirty states. Refactor the old blf dirty check into a helper function to be used by the assert. Signed-off-by:
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Brian Foster authored
commit a4f6cf6b upstream. It checks a single flag and has one caller. It probably isn't worth its own function. Signed-off-by:
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Omar Sandoval authored
commit f2e9ad21 upstream. After xfs_ifree_cluster() finds an inode in the radix tree and verifies that the inode number is what it expected, xfs_reclaim_inode() can swoop in and free it. xfs_ifree_cluster() will then happily continue working on the freed inode. Most importantly, it will mark the inode stale, which will probably be overwritten when the inode slab object is reallocated, but if it has already been reallocated then we can end up with an inode spuriously marked stale. In 8a17d7dd ("xfs: mark reclaimed inodes invalid earlier") we added a second check to xfs_iflush_cluster() to detect this race, but the similar RCU lookup in xfs_ifree_cluster() needs the same treatment. Signed-off-by:
Omar Sandoval <osandov@fb.com> Reviewed-by:
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Darrick J. Wong authored
commit 799ea9e9 upstream. When we introduced the bmap redo log items, we set MS_ACTIVE on the mountpoint and XFS_IRECOVERY on the inode to prevent unlinked inodes from being truncated prematurely during log recovery. This also had the effect of putting linked inodes on the lru instead of evicting them. Unfortunately, we neglected to find all those unreferenced lru inodes and evict them after finishing log recovery, which means that we leak them if anything goes wrong in the rest of xfs_mountfs, because the lru is only cleaned out on unmount. Therefore, evict unreferenced inodes in the lru list immediately after clearing MS_ACTIVE. Fixes: 17c12bcd ("xfs: when replaying bmap operations, don't let unlinked inodes get reaped") Signed-off-by:
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Carlos Maiolino authored
commit 2d32311c upstream. In a filesystem without finobt, the Space manager selects an AG to alloc a new inode, where xfs_dialloc_ag_inobt() will search the AG for the free slot chunk. When the new inode is in the same AG as its parent, the btree will be searched starting on the parent's record, and then retried from the top if no slot is available beyond the parent's record. To exit this loop though, xfs_dialloc_ag_inobt() relies on the fact that the btree must have a free slot available, once its callers relied on the agi->freecount when deciding how/where to allocate this new inode. In the case when the agi->freecount is corrupted, showing available inodes in an AG, when in fact there is none, this becomes an infinite loop. Add a way to stop the loop when a free slot is not found in the btree, making the function to fall into the whole AG scan which will then, be able to detect the corruption and shut the filesystem down. As pointed by Brian, this might impact performance, giving the fact we don't reset the search distance anymore when we reach the end of the tree, giving it fewer tries before falling back to the whole AG search, but it will only affect searches that start within 10 records to the end of the tree. Signed-off-by:
Carlos Maiolino <cmaiolino@redhat.com> Reviewed-by:
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Brian Foster authored
commit a4c9b34d upstream. Torn write and tail overwrite detection both trigger only on -EFSBADCRC errors. While this is the most likely failure scenario for each condition, -EFSCORRUPTED is still possible in certain cases depending on what ends up on disk when a torn write or partial tail overwrite occurs. For example, an invalid log record h_len can lead to an -EFSCORRUPTED error when running the log recovery CRC pass. Therefore, update log head and tail verification to trigger the associated head/tail fixups in the event of -EFSCORRUPTED errors along with -EFSBADCRC. Also, -EFSCORRUPTED can currently be returned from xlog_do_recovery_pass() before rhead_blk is initialized if the first record encountered happens to be corrupted. This leads to an incorrect 'first_bad' return value. Initialize rhead_blk earlier in the function to address that problem as well. Signed-off-by:
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Brian Foster authored
commit 4a4f66ea upstream. If we consider the case where the tail (T) of the log is pinned long enough for the head (H) to push and block behind the tail, we can end up blocked in the following state without enough free space (f) in the log to satisfy a transaction reservation: 0 phys. log N [-------HffT---H'--T'---] The last good record in the log (before H) refers to T. The tail eventually pushes forward (T') leaving more free space in the log for writes to H. At this point, suppose space frees up in the log for the maximum of 8 in-core log buffers to start flushing out to the log. If this pushes the head from H to H', these next writes overwrite the previous tail T. This is safe because the items logged from T to T' have been written back and removed from the AIL. If the next log writes (H -> H') happen to fail and result in partial records in the log, the filesystem shuts down having overwritten T with invalid data. Log recovery correctly locates H on the subsequent mount, but H still refers to the now corrupted tail T. This results in log corruption errors and recovery failure. Since the tail overwrite results from otherwise correct runtime behavior, it is up to log recovery to try and deal with this situation. Update log recovery tail verification to run a CRC pass from the first record past the tail to the head. This facilitates error detection at T and moves the recovery tail to the first good record past H' (similar to truncating the head on torn write detection). If corruption is detected beyond the range possibly affected by the max number of iclogs, the log is legitimately corrupted and log recovery failure is expected. Signed-off-by:
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Brian Foster authored
commit 5297ac1f upstream. Log tail verification currently only occurs when torn writes are detected at the head of the log. This was introduced because a change in the head block due to torn writes can lead to a change in the tail block (each log record header references the current tail) and the tail block should be verified before log recovery proceeds. Tail corruption is possible outside of torn write scenarios, however. For example, partial log writes can be detected and cleared during the initial head/tail block discovery process. If the partial write coincides with a tail overwrite, the log tail is corrupted and recovery fails. To facilitate correct handling of log tail overwites, update log recovery to always perform tail verification. This is necessary to detect potential tail overwrite conditions when torn writes may not have occurred. This changes normal (i.e., no torn writes) recovery behavior slightly to detect and return CRC related errors near the tail before actual recovery starts. Signed-off-by:
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Brian Foster authored
commit 284f1c2c upstream. The high-level log recovery algorithm consists of two loops that walk the physical log and process log records from the tail to the head. The first loop handles the case where the tail is beyond the head and processes records up to the end of the physical log. The subsequent loop processes records from the beginning of the physical log to the head. Because log records can wrap around the end of the physical log, the first loop mentioned above must handle this case appropriately. Records are processed from in-core buffers, which means that this algorithm must split the reads of such records into two partial I/Os: 1.) from the beginning of the record to the end of the log and 2.) from the beginning of the log to the end of the record. This is further complicated by the fact that the log record header and log record data are read into independent buffers. The current handling of each buffer correctly splits the reads when either the header or data starts before the end of the log and wraps around the end. The data read does not correctly handle the case where the prior header read wrapped or ends on the physical log end boundary. blk_no is incremented to or beyond the log end after the header read to point to the record data, but the split data read logic triggers, attempts to read from an invalid log block and ultimately causes log recovery to fail. This can be reproduced fairly reliably via xfstests tests generic/047 and generic/388 with large iclog sizes (256k) and small (10M) logs. If the record header read has pushed beyond the end of the physical log, the subsequent data read is actually contiguous. Update the data read logic to detect the case where blk_no has wrapped, mod it against the log size to read from the correct address and issue one contiguous read for the log data buffer. The log record is processed as normal from the buffer(s), the loop exits after the current iteration and the subsequent loop picks up with the first new record after the start of the log. Signed-off-by:
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Carlos Maiolino authored
commit d3a304b6 upstream. When a buffer has been failed during writeback, the inode items into it are kept flush locked, and are never resubmitted due the flush lock, so, if any buffer fails to be written, the items in AIL are never written to disk and never unlocked. This causes unmount operation to hang due these items flush locked in AIL, but this also causes the items in AIL to never be written back, even when the IO device comes back to normal. I've been testing this patch with a DM-thin device, creating a filesystem larger than the real device. When writing enough data to fill the DM-thin device, XFS receives ENOSPC errors from the device, and keep spinning on xfsaild (when 'retry forever' configuration is set). At this point, the filesystem can not be unmounted because of the flush locked items in AIL, but worse, the items in AIL are never retried at all (once xfs_inode_item_push() will skip the items that are flush locked), even if the underlying DM-thin device is expanded to the proper size. This patch fixes both cases, retrying any item that has been failed previously, using the infra-structure provided by the previous patch. Reviewed-by:
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Carlos Maiolino authored
commit 0b80ae6e upstream. With the current code, XFS never re-submit a failed buffer for IO, because the failed item in the buffer is kept in the flush locked state forever. To be able to resubmit an log item for IO, we need a way to mark an item as failed, if, for any reason the buffer which the item belonged to failed during writeback. Add a new log item callback to be used after an IO completion failure and make the needed clean ups. Reviewed-by:
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Eric Sandeen authored
commit 6f4a1eef upstream. When we do log recovery on a readonly mount, unlinked inode processing does not happen due to the readonly checks in xfs_inactive(), which are trying to prevent any I/O on a readonly mount. This is misguided - we do I/O on readonly mounts all the time, for consistency; for example, log recovery. So do the same RDONLY flag twiddling around xfs_log_mount_finish() as we do around xfs_log_mount(), for the same reason. This all cries out for a big rework but for now this is a simple fix to an obvious problem. Signed-off-by:
Eric Sandeen <sandeen@redhat.com> Reviewed-by:
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Eric Sandeen authored
commit 757a69ef upstream. There are dueling comments in the xfs code about intent for log writes when unmounting a readonly filesystem. In xfs_mountfs, we see the intent: /* * Now the log is fully replayed, we can transition to full read-only * mode for read-only mounts. This will sync all the metadata and clean * the log so that the recovery we just performed does not have to be * replayed again on the next mount. */ and it calls xfs_quiesce_attr(), but by the time we get to xfs_log_unmount_write(), it returns early for a RDONLY mount: * Don't write out unmount record on read-only mounts. Because of this, sequential ro mounts of a filesystem with a dirty log will replay the log each time, which seems odd. Fix this by writing an unmount record even for RO mounts, as long as norecovery wasn't specified (don't write a clean log record if a dirty log may still be there!) and the log device is writable. Signed-off-by:
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Dan Williams authored
commit 58738c49 upstream. Dan reports: The patch 62232e45: "libnvdimm: control (ioctl) messages for nvdimm_bus and nvdimm devices" from Jun 8, 2015, leads to the following static checker warning: drivers/nvdimm/bus.c:1018 __nd_ioctl() warn: integer overflows 'buf_len' From a casual review, this seems like it might be a real bug. On the first iteration we load some data into in_env[]. On the second iteration we read a use controlled "in_size" from nd_cmd_in_size(). It can go up to UINT_MAX - 1. A high number means we will fill the whole in_env[] buffer. But we potentially keep looping and adding more to in_len so now it can be any value. It simple enough to change, but it feels weird that we keep looping even though in_env is totally full. Shouldn't we just return an error if we don't have space for desc->in_num. We keep looping because the size of the total input is allowed to be bigger than the 'envelope' which is a subset of the payload that tells us how much data to expect. For safety explicitly check that buf_len does not overflow which is what the checker flagged. Fixes: 62232e45: "libnvdimm: control (ioctl) messages for nvdimm_bus..." Reported-by:
Dan Carpenter <dan.carpenter@oracle.com> Signed-off-by:
Dan Williams <dan.j.williams@intel.com> Signed-off-by:
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Christophe Jaillet authored
commit ed36b4db upstream. Check memory allocation failures and return -ENOMEM in such cases, as already done few lines below for another memory allocation. This avoids NULL pointers dereference. Fixes: 14e49454 ("libnvdimm, btt: BTT updates for UEFI 2.7 format") Signed-off-by:
Christophe JAILLET <christophe.jaillet@wanadoo.fr> Reviewed-by:
Vishal Verma <vishal.l.verma@intel.com> Signed-off-by:
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Eric Biggers authored
commit a47f68d6 upstream. IDR only supports non-negative IDs. There used to be a 'WARN_ON_ONCE(id < 0)' in idr_replace(), but it was intentionally removed by commit 2e1c9b28 ("idr: remove WARN_ON_ONCE() on negative IDs"). Then it was added back by commit 0a835c4f ("Reimplement IDR and IDA using the radix tree"). However it seems that adding it back was a mistake, given that some users such as drm_gem_handle_delete() (DRM_IOCTL_GEM_CLOSE) pass in a value from userspace to idr_replace(), allowing the WARN_ON_ONCE to be triggered. drm_gem_handle_delete() actually just wants idr_replace() to return an error code if the ID is not allocated, including in the case where the ID is invalid (negative). So once again remove the bogus WARN_ON_ONCE(). This bug was found by syzkaller, which encountered the following warning: WARNING: CPU: 3 PID: 3008 at lib/idr.c:157 idr_replace+0x1d8/0x240 lib/idr.c:157 Kernel panic - not syncing: panic_on_warn set ... CPU: 3 PID: 3008 Comm: syzkaller218828 Not tainted 4.13.0-rc4-next-20170811 #2 Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS Bochs 01/01/2011 Call Trace: fixup_bug+0x40/0x90 arch/x86/kernel/traps.c:190 do_trap_no_signal arch/x86/kernel/traps.c:224 [inline] do_trap+0x260/0x390 arch/x86/kernel/traps.c:273 do_error_trap+0x120/0x390 arch/x86/kernel/traps.c:310 do_invalid_op+0x1b/0x20 arch/x86/kernel/traps.c:323 invalid_op+0x1e/0x30 arch/x86/entry/entry_64.S:930 RIP: 0010:idr_replace+0x1d8/0x240 lib/idr.c:157 RSP: 0018:ffff8800394bf9f8 EFLAGS: 00010297 RAX: ffff88003c6c60c0 RBX: 1ffff10007297f43 RCX: 0000000000000000 RDX: 0000000000000000 RSI: 0000000000000000 RDI: ffff8800394bfa78 RBP: ffff8800394bfae0 R08: ffffffff82856487 R09: 0000000000000000 R10: ffff8800394bf9a8 R11: ffff88006c8bae28 R12: ffffffffffffffff R13: ffff8800394bfab8 R14: dffffc0000000000 R15: ffff8800394bfbc8 drm_gem_handle_delete+0x33/0xa0 drivers/gpu/drm/drm_gem.c:297 drm_gem_close_ioctl+0xa1/0xe0 drivers/gpu/drm/drm_gem.c:671 drm_ioctl_kernel+0x1e7/0x2e0 drivers/gpu/drm/drm_ioctl.c:729 drm_ioctl+0x72e/0xa50 drivers/gpu/drm/drm_ioctl.c:825 vfs_ioctl fs/ioctl.c:45 [inline] do_vfs_ioctl+0x1b1/0x1520 fs/ioctl.c:685 SYSC_ioctl fs/ioctl.c:700 [inline] SyS_ioctl+0x8f/0xc0 fs/ioctl.c:691 entry_SYSCALL_64_fastpath+0x1f/0xbe Here is a C reproducer: #include <fcntl.h> #include <stddef.h> #include <stdint.h> #include <sys/ioctl.h> #include <drm/drm.h> int main(void) { int cardfd = open("/dev/dri/card0", O_RDONLY); ioctl(cardfd, DRM_IOCTL_GEM_CLOSE, &(struct drm_gem_close) { .handle = -1 } ); } Link: http://lkml.kernel.org/r/20170906235306.20534-1-ebiggers3@gmail.com Fixes: 0a835c4f ("Reimplement IDR and IDA using the radix tree") Signed-off-by:
Eric Biggers <ebiggers@google.com> Acked-by:
Tejun Heo <tj@kernel.org> Cc: Dmitry Vyukov <dvyukov@google.com> Cc: Matthew Wilcox <mawilcox@microsoft.com> Signed-off-by:
Andrew Morton <akpm@linux-foundation.org> Signed-off-by:
Linus Torvalds <torvalds@linux-foundation.org> Signed-off-by:
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Miklos Szeredi authored
commit 5d6d3a30 upstream. Commit 0b6e9ea0 ("fuse: Add support for pid namespaces") broke Sandstorm.io development tools, which have been sending FUSE file descriptors across PID namespace boundaries since early 2014. The above patch added a check that prevented I/O on the fuse device file descriptor if the pid namespace of the reader/writer was different from the pid namespace of the mounter. With this change passing the device file descriptor to a different pid namespace simply doesn't work. The check was added because pids are transferred to/from the fuse userspace server in the namespace registered at mount time. To fix this regression, remove the checks and do the following: 1) the pid in the request header (the pid of the task that initiated the filesystem operation) is translated to the reader's pid namespace. If a mapping doesn't exist for this pid, then a zero pid is used. Note: even if a mapping would exist between the initiator task's pid namespace and the reader's pid namespace the pid will be zero if either mapping from initator's to mounter's namespace or mapping from mounter's to reader's namespace doesn't exist. 2) The lk.pid value in setlk/setlkw requests and getlk reply is left alone. Userspace should not interpret this value anyway. Also allow the setlk/setlkw operations if the pid of the task cannot be represented in the mounter's namespace (pid being zero in that case). Reported-by:
Kenton Varda <kenton@sandstorm.io> Signed-off-by:
Miklos Szeredi <mszeredi@redhat.com> Fixes: 0b6e9ea0 ("fuse: Add support for pid namespaces") Cc: Eric W. Biederman <ebiederm@xmission.com> Cc: Seth Forshee <seth.forshee@canonical.com> Signed-off-by:
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Amir Goldstein authored
commit 939ae4ef upstream. Commit b9ac5c27 ("ovl: hash overlay non-dir inodes by copy up origin") verifies that the origin lower inode stored in the overlayfs inode matched the inode of a copy up origin dentry found by lookup. There is a false positive result in that check when lower fs does not support file handles and copy up origin cannot be followed by file handle at lookup time. The false negative happens when finding an overlay inode in cache on a copied up overlay dentry lookup. The overlay inode still 'remembers' the copy up origin inode, but the copy up origin dentry is not available for verification. Relax the check in case copy up origin dentry is not available. Fixes: b9ac5c27 ("ovl: hash overlay non-dir inodes by copy up...") Reported-by:
Jordi Pujol <jordipujolp@gmail.com> Signed-off-by:
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Tony Luck authored
commit ce0fa3e5 upstream. Speculative processor accesses may reference any memory that has a valid page table entry. While a speculative access won't generate a machine check, it will log the error in a machine check bank. That could cause escalation of a subsequent error since the overflow bit will be then set in the machine check bank status register. Code has to be double-plus-tricky to avoid mentioning the 1:1 virtual address of the page we want to map out otherwise we may trigger the very problem we are trying to avoid. We use a non-canonical address that passes through the usual Linux table walking code to get to the same "pte". Thanks to Dave Hansen for reviewing several iterations of this. Also see: http://marc.info/?l=linux-mm&m=149860136413338&w=2Signed-off-by:
Tony Luck <tony.luck@intel.com> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Andy Lutomirski <luto@kernel.org> Cc: Borislav Petkov <bp@alien8.de> Cc: Borislav Petkov <bp@suse.de> Cc: Brian Gerst <brgerst@gmail.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: Denys Vlasenko <dvlasenk@redhat.com> Cc: Elliott, Robert (Persistent Memory) <elliott@hpe.com> Cc: H. Peter Anvin <hpa@zytor.com> Cc: Josh Poimboeuf <jpoimboe@redhat.com> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: linux-mm@kvack.org Link: http://lkml.kernel.org/r/20170816171803.28342-1-tony.luck@intel.comSigned-off-by:
Ingo Molnar <mingo@kernel.org> Signed-off-by:
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Andy Lutomirski authored
commit e137a4d8 upstream. Switching FS and GS is a mess, and the current code is still subtly wrong: it assumes that "Loading a nonzero value into FS sets the index and base", which is false on AMD CPUs if the value being loaded is 1, 2, or 3. (The current code came from commit 3e2b68d7 ("x86/asm, sched/x86: Rewrite the FS and GS context switch code"), which made it better but didn't fully fix it.) Rewrite it to be much simpler and more obviously correct. This should fix it fully on AMD CPUs and shouldn't adversely affect performance. Signed-off-by:
Andy Lutomirski <luto@kernel.org> Cc: Borislav Petkov <bp@alien8.de> Cc: Borislav Petkov <bpetkov@suse.de> Cc: Brian Gerst <brgerst@gmail.com> Cc: Chang Seok <chang.seok.bae@intel.com> Cc: Denys Vlasenko <dvlasenk@redhat.com> Cc: H. Peter Anvin <hpa@zytor.com> Cc: Josh Poimboeuf <jpoimboe@redhat.com> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Signed-off-by:
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Andy Lutomirski authored
commit 9584d98b upstream. In ELF_COPY_CORE_REGS, we're copying from the current task, so accessing thread.fsbase and thread.gsbase makes no sense. Just read the values from the CPU registers. In practice, the old code would have been correct most of the time simply because thread.fsbase and thread.gsbase usually matched the CPU registers. Signed-off-by:
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Andy Lutomirski authored
commit 767d035d upstream. execve used to leak FSBASE and GSBASE on AMD CPUs. Fix it. The security impact of this bug is small but not quite zero -- it could weaken ASLR when a privileged task execs a less privileged program, but only if program changed bitness across the exec, or the child binary was highly unusual or actively malicious. A child program that was compromised after the exec would not have access to the leaked base. Signed-off-by:
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Bernat, Yehezkel authored
commit e545f0d8 upstream. If secure authentication of a devices fails, either because the device already has another key uploaded, or there is some other error sending challenge to the device, and the user only wants to approve the device just once (without a new key being uploaded to the device) the current implementation does not allow this because the key cannot be cleared once set even if we allow it to be changed. Make this scenario possible and allow clearing the key by writing empty string to the key sysfs file. Signed-off-by:
Yehezkel Bernat <yehezkel.bernat@intel.com> Acked-by:
Mika Westerberg <mika.westerberg@linux.intel.com> Signed-off-by:
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Bernat, Yehezkel authored
commit 0956e411 upstream. Non-root user may read the key back after root wrote it there. This removes read access to everyone but root. Signed-off-by:
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