| Commit message (Collapse) | Author | Age |
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The stack size was hard coded to 0x2000, use the standard THREAD_SIZE
definition loaded from thread_info.h.
Signed-off-by: Stefan Kristiansson <stefan.kristiansson@saunalahti.fi>
[shorne@gmail.com: Added body to the commit message]
Signed-off-by: Stafford Horne <shorne@gmail.com>
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By slightly reorganizing the code, the number of registers
used in the tlb miss handlers can be reduced by two,
thus removing the need to save them to memory.
Also, some dead and commented out code is removed.
No functional change.
Signed-off-by: Stefan Kristiansson <stefan.kristiansson@saunalahti.fi>
Signed-off-by: Jonas Bonn <jonas@southpole.se>
Signed-off-by: Stafford Horne <shorne@gmail.com>
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Motivation for this is to be able to print the way information
properly in print_cpuinfo(), instead of hardcoding it to one.
Signed-off-by: Stefan Kristiansson <stefan.kristiansson@saunalahti.fi>
Signed-off-by: Jonas Bonn <jonas@southpole.se>
[shorne@gmail.com fixed conflict with show_cpuinfo change]
Signed-off-by: Stafford Horne <shorne@gmail.com>
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The sparse IRQ framework is preferred nowadays so switch over to it.
Signed-off-by: Jonas Bonn <jonas@southpole.se>
Signed-off-by: Stafford Horne <shorne@gmail.com>
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Pull sparc updates from David Miller:
1) Support multiple huge page sizes, from Nitin Gupta.
2) Improve boot time on large memory configurations, from Pavel
Tatashin.
3) Make BRK handling more consistent and documented, from Vijay Kumar.
* git://git.kernel.org/pub/scm/linux/kernel/git/davem/sparc:
sparc64: Fix build error in flush_tsb_user_page
sparc64: memblock resizes are not handled properly
sparc64: use latency groups to improve add_node_ranges speed
sparc64: Add 64K page size support
sparc64: Multi-page size support
Documentation/sparc: Steps for sending break on sunhv console
sparc64: Send break twice from console to return to boot prom
sparc64: Migrate hvcons irq to panicked cpu
sparc64: Set cpu state to offline when stopped
sunvdc: Add support for setting physical sector size
sparc64: fix for user probes in high memory
sparc: topology_64.h: Fix condition for including cpudata.h
sparc32: mm: srmmu: add __ro_after_init to sparc32_cachetlb_ops structures
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Patch "sparc64: Add 64K page size support"
unconditionally used __flush_huge_tsb_one_entry()
which is available only when hugetlb support is
enabled.
Another issue was incorrect TSB flushing for 64K
pages in flush_tsb_user().
Signed-off-by: Nitin Gupta <nitin.m.gupta@oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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In add_node_ranges() when memblock resize happens, the iterator keeps using
the previous freed array. This bug cause hangs on machine where there are
over 128 memory blocks during boot. For example, on machines where memory
interleaving is small.
The problem is seen on T4-4 because it cant have 2T of memory, and memory
is interleaved at 8G. So we have 2T/8G = 256 regions to set node IDs. The
starting size of regions array is 128. Thus, we have to double at least one
time (actually we have to double twice because some memory is already
reserved and thus we need more than 256 regions). We start using an
incorrect pointer to the array after the first doubling.
Signed-off-by: Pavel Tatashin <pasha.tatashin@oracle.com>
Signed-off-by: Babu Moger <babu.moger@oracle.com>
Reviewed-by: Babu Moger <babu.moger@oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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add_node_ranges() takes 2.6s - 3.6s per 1T of boot time. On machine with 6T
memory it takes 15.4s, on 32T it would take 82s-115s of boot time.
This function sets NUMA ids for memory blocks, and scans the whole memory a
page at a time to do so. But, we could use values in latency groups mask
and match to determine the boundaries without checking every single page.
With the fix the add_node_ranges() time is reduced from 15.4s down to 0.2s
on machine with 6T memory.
Signed-off-by: Pavel Tatashin <pasha.tatashin@oracle.com>
Reviewed-by: Babu Moger <babu.moger@oracle.com>
Reviewed-by: Bob Picco <bob.picco@oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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This patch depends on:
[v6] sparc64: Multi-page size support
- Testing
Tested on Sonoma by running stream benchmark instance which allocated
48G worth of 64K pages.
boot params: default_hugepagesz=64K hugepagesz=64K hugepages=1310720
Signed-off-by: Nitin Gupta <nitin.m.gupta@oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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Add support for using multiple hugepage sizes simultaneously
on mainline. Currently, support for 256M has been added which
can be used along with 8M pages.
Page tables are set like this (e.g. for 256M page):
VA + (8M * x) -> PA + (8M * x) (sz bit = 256M) where x in [0, 31]
and TSB is set similarly:
VA + (4M * x) -> PA + (4M * x) (sz bit = 256M) where x in [0, 63]
- Testing
Tested on Sonoma (which supports 256M pages) by running stream
benchmark instances in parallel: one instance uses 8M pages and
another uses 256M pages, consuming 48G each.
Boot params used:
default_hugepagesz=256M hugepagesz=256M hugepages=300 hugepagesz=8M
hugepages=10000
Signed-off-by: Nitin Gupta <nitin.m.gupta@oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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Vijay Kumar says:
====================
sparc64: Jump to boot prom from console on panic
V3 changes:
- patch 02/04: Added SERIAL_SUNHV conditional group for
sunhv_migrate_hvcons_irq in smp_send_stop().
V2 changes:
- Added cover letter patch
Currently Stop-A (L1A) does not make the kernel switch to OBP on panic. This
patchset addresses this issue. Also, now we can cause a jump to OBP by sending
'break' twice from sunhv console. On bare metal, one can send a break by
typing Esc + 'B' + Sysrq (or whatever). On LDOM, press Ctrl + ] in telnet,
and then "send break" at the telnet prompt.
====================
Signed-off-by: David S. Miller <davem@davemloft.net>
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Documented the steps for sending break on sunhv console.
Signed-off-by: Vijay Kumar <vijay.ac.kumar@oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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Now we can also jump to boot prom from sunhv console by sending
break twice on console for both running and panicked kernel
cases.
Signed-off-by: Vijay Kumar <vijay.ac.kumar@oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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On panic, all other CPUs are stopped except the one which had
hit panic. To keep console alive, we need to migrate hvcons irq
to panicked CPU.
Signed-off-by: Vijay Kumar <vijay.ac.kumar@oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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CPU needs to be marked offline before stopping it. When not marked
offline, the xcall receives HV_EWOULDBLOCK and so assumes that not all
CPUs received the message, and retries. After 10000 retries, it finally
fails with fatal mondo timeout.
Signed-off-by: Vijay Kumar <vijay.ac.kumar@oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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Physical sector size is supported in v1.2 of the vDisk protocol and
should be set if available. If protocol version 1.2 is used and the
physical disk size is unavailable, then the disk is considered busy.
Signed-off-by: Liam R. Howlett <Liam.Howlett@Oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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When returning from the user probe code into userspace process, PC & NPC are
truncated to 32 bits.
Due to shared libraries getting loaded very high in the virtual address
space of
the process, placing a user probe inside a shared library makes the kernel
return into the process at the wrong address, causing it to seg'fault
most of
the time.
This patch prevents truncating PC and NPC.
Signed-off-by: Eric Saint Etienne <eric.saint.etienne@oracle.com>
Reviewed-by: David Aldridge <david.j.aldridge@oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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We currently define macros referring to cpu_data if CONFIG_SMP is
defined, but only include the declaration if CONFIG_NUMA is defined.
Fixes: 541cc39433a8 ("sparc: fix a building error reported by kbuild")
Signed-off-by: Ben Hutchings <ben@decadent.org.uk>
Acked-by: Gonglei <arei.gonglei@huawei.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
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The objects viking_ops, viking_sun4d_smp_ops and smp_cachetlb_ops of
type sparc32_cachetlb_ops are not modified anywhere after getting modified
in the init functions. Inside init their reference is also stored in a
pointer of type const struct sparc32_cachetlb_ops *. So these structures
are never modified after init, therefore add __ro_after to the declaration
of these structures.
Signed-off-by: Bhumika Goyal <bhumirks@gmail.com>
Reviewed-by: Kees Cook <keescook@chromium.org>
Signed-off-by: David S. Miller <davem@davemloft.net>
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Pull md updates from Shaohua Li:
"Mainly fixes bugs and improves performance:
- Improve scalability for raid1 from Coly
- Improve raid5-cache read performance, disk efficiency and IO
pattern from Song and me
- Fix a race condition of disk hotplug for linear from Coly
- A few cleanup patches from Ming and Byungchul
- Fix a memory leak from Neil
- Fix WRITE SAME IO failure from me
- Add doc for raid5-cache from me"
* 'for-next' of git://git.kernel.org/pub/scm/linux/kernel/git/shli/md: (23 commits)
md/raid1: fix write behind issues introduced by bio_clone_bioset_partial
md/raid1: handle flush request correctly
md/linear: shutup lockdep warnning
md/raid1: fix a use-after-free bug
RAID1: avoid unnecessary spin locks in I/O barrier code
RAID1: a new I/O barrier implementation to remove resync window
md/raid5: Don't reinvent the wheel but use existing llist API
md: fast clone bio in bio_clone_mddev()
md: remove unnecessary check on mddev
md/raid1: use bio_clone_bioset_partial() in case of write behind
md: fail if mddev->bio_set can't be created
block: introduce bio_clone_bioset_partial()
md: disable WRITE SAME if it fails in underlayer disks
md/raid5-cache: exclude reclaiming stripes in reclaim check
md/raid5-cache: stripe reclaim only counts valid stripes
MD: add doc for raid5-cache
Documentation: move MD related doc into a separate dir
md: ensure md devices are freed before module is unloaded.
md/r5cache: improve journal device efficiency
md/r5cache: enable chunk_aligned_read with write back cache
...
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There are two issues, introduced by commit 8e58e32(md/raid1: use
bio_clone_bioset_partial() in case of write behind):
- bio_clone_bioset_partial() uses bytes instead of sectors as parameters
- in writebehind mode, we return bio if all !writemostly disk bios finish,
which could happen before writemostly disk bios run. So all
writemostly disk bios should have their bvec. Here we just make sure
all bios are cloned instead of fast cloned.
Reviewed-by: Ming Lei <tom.leiming@gmail.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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I got a warning triggered in align_to_barrier_unit_end. It's a flush
request so sectors == 0. The flush request happens to work well without
the new barrier patch, but we'd better handle it explictly.
Cc: NeilBrown <neilb@suse.com>
Acked-by: Coly Li <colyli@suse.de>
Signed-off-by: Shaohua Li <shli@fb.com>
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Commit 03a9e24(md linear: fix a race between linear_add() and
linear_congested()) introduces the warnning.
Acked-by: Coly Li <colyli@suse.de>
Signed-off-by: Shaohua Li <shli@fb.com>
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Commit fd76863 (RAID1: a new I/O barrier implementation to remove resync
window) introduces a user-after-free bug.
Signed-off-by: Shaohua Li <shli@fb.com>
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When I run a parallel reading performan testing on a md raid1 device with
two NVMe SSDs, I observe very bad throughput in supprise: by fio with 64KB
block size, 40 seq read I/O jobs, 128 iodepth, overall throughput is
only 2.7GB/s, this is around 50% of the idea performance number.
The perf reports locking contention happens at allow_barrier() and
wait_barrier() code,
- 41.41% fio [kernel.kallsyms] [k] _raw_spin_lock_irqsave
- _raw_spin_lock_irqsave
+ 89.92% allow_barrier
+ 9.34% __wake_up
- 37.30% fio [kernel.kallsyms] [k] _raw_spin_lock_irq
- _raw_spin_lock_irq
- 100.00% wait_barrier
The reason is, in these I/O barrier related functions,
- raise_barrier()
- lower_barrier()
- wait_barrier()
- allow_barrier()
They always hold conf->resync_lock firstly, even there are only regular
reading I/Os and no resync I/O at all. This is a huge performance penalty.
The solution is a lockless-like algorithm in I/O barrier code, and only
holding conf->resync_lock when it has to.
The original idea is from Hannes Reinecke, and Neil Brown provides
comments to improve it. I continue to work on it, and make the patch into
current form.
In the new simpler raid1 I/O barrier implementation, there are two
wait barrier functions,
- wait_barrier()
Which calls _wait_barrier(), is used for regular write I/O. If there is
resync I/O happening on the same I/O barrier bucket, or the whole
array is frozen, task will wait until no barrier on same barrier bucket,
or the whold array is unfreezed.
- wait_read_barrier()
Since regular read I/O won't interfere with resync I/O (read_balance()
will make sure only uptodate data will be read out), it is unnecessary
to wait for barrier in regular read I/Os, waiting in only necessary
when the whole array is frozen.
The operations on conf->nr_pending[idx], conf->nr_waiting[idx], conf->
barrier[idx] are very carefully designed in raise_barrier(),
lower_barrier(), _wait_barrier() and wait_read_barrier(), in order to
avoid unnecessary spin locks in these functions. Once conf->
nr_pengding[idx] is increased, a resync I/O with same barrier bucket index
has to wait in raise_barrier(). Then in _wait_barrier() if no barrier
raised in same barrier bucket index and array is not frozen, the regular
I/O doesn't need to hold conf->resync_lock, it can just increase
conf->nr_pending[idx], and return to its caller. wait_read_barrier() is
very similar to _wait_barrier(), the only difference is it only waits when
array is frozen. For heavy parallel reading I/Os, the lockless I/O barrier
code almostly gets rid of all spin lock cost.
This patch significantly improves raid1 reading peroformance. From my
testing, a raid1 device built by two NVMe SSD, runs fio with 64KB
blocksize, 40 seq read I/O jobs, 128 iodepth, overall throughput
increases from 2.7GB/s to 4.6GB/s (+70%).
Changelog
V4:
- Change conf->nr_queued[] to atomic_t.
- Define BARRIER_BUCKETS_NR_BITS by (PAGE_SHIFT - ilog2(sizeof(atomic_t)))
V3:
- Add smp_mb__after_atomic() as Shaohua and Neil suggested.
- Change conf->nr_queued[] from atomic_t to int.
- Change conf->array_frozen from atomic_t back to int, and use
READ_ONCE(conf->array_frozen) to check value of conf->array_frozen
in _wait_barrier() and wait_read_barrier().
- In _wait_barrier() and wait_read_barrier(), add a call to
wake_up(&conf->wait_barrier) after atomic_dec(&conf->nr_pending[idx]),
to fix a deadlock between _wait_barrier()/wait_read_barrier and
freeze_array().
V2:
- Remove a spin_lock/unlock pair in raid1d().
- Add more code comments to explain why there is no racy when checking two
atomic_t variables at same time.
V1:
- Original RFC patch for comments.
Signed-off-by: Coly Li <colyli@suse.de>
Cc: Shaohua Li <shli@fb.com>
Cc: Hannes Reinecke <hare@suse.com>
Cc: Johannes Thumshirn <jthumshirn@suse.de>
Cc: Guoqing Jiang <gqjiang@suse.com>
Reviewed-by: Neil Brown <neilb@suse.de>
Signed-off-by: Shaohua Li <shli@fb.com>
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'Commit 79ef3a8aa1cb ("raid1: Rewrite the implementation of iobarrier.")'
introduces a sliding resync window for raid1 I/O barrier, this idea limits
I/O barriers to happen only inside a slidingresync window, for regular
I/Os out of this resync window they don't need to wait for barrier any
more. On large raid1 device, it helps a lot to improve parallel writing
I/O throughput when there are background resync I/Os performing at
same time.
The idea of sliding resync widow is awesome, but code complexity is a
challenge. Sliding resync window requires several variables to work
collectively, this is complexed and very hard to make it work correctly.
Just grep "Fixes: 79ef3a8aa1" in kernel git log, there are 8 more patches
to fix the original resync window patch. This is not the end, any further
related modification may easily introduce more regreassion.
Therefore I decide to implement a much simpler raid1 I/O barrier, by
removing resync window code, I believe life will be much easier.
The brief idea of the simpler barrier is,
- Do not maintain a global unique resync window
- Use multiple hash buckets to reduce I/O barrier conflicts, regular
I/O only has to wait for a resync I/O when both them have same barrier
bucket index, vice versa.
- I/O barrier can be reduced to an acceptable number if there are enough
barrier buckets
Here I explain how the barrier buckets are designed,
- BARRIER_UNIT_SECTOR_SIZE
The whole LBA address space of a raid1 device is divided into multiple
barrier units, by the size of BARRIER_UNIT_SECTOR_SIZE.
Bio requests won't go across border of barrier unit size, that means
maximum bio size is BARRIER_UNIT_SECTOR_SIZE<<9 (64MB) in bytes.
For random I/O 64MB is large enough for both read and write requests,
for sequential I/O considering underlying block layer may merge them
into larger requests, 64MB is still good enough.
Neil also points out that for resync operation, "we want the resync to
move from region to region fairly quickly so that the slowness caused
by having to synchronize with the resync is averaged out over a fairly
small time frame". For full speed resync, 64MB should take less then 1
second. When resync is competing with other I/O, it could take up a few
minutes. Therefore 64MB size is fairly good range for resync.
- BARRIER_BUCKETS_NR
There are BARRIER_BUCKETS_NR buckets in total, which is defined by,
#define BARRIER_BUCKETS_NR_BITS (PAGE_SHIFT - 2)
#define BARRIER_BUCKETS_NR (1<<BARRIER_BUCKETS_NR_BITS)
this patch makes the bellowed members of struct r1conf from integer
to array of integers,
- int nr_pending;
- int nr_waiting;
- int nr_queued;
- int barrier;
+ int *nr_pending;
+ int *nr_waiting;
+ int *nr_queued;
+ int *barrier;
number of the array elements is defined as BARRIER_BUCKETS_NR. For 4KB
kernel space page size, (PAGE_SHIFT - 2) indecates there are 1024 I/O
barrier buckets, and each array of integers occupies single memory page.
1024 means for a request which is smaller than the I/O barrier unit size
has ~0.1% chance to wait for resync to pause, which is quite a small
enough fraction. Also requesting single memory page is more friendly to
kernel page allocator than larger memory size.
- I/O barrier bucket is indexed by bio start sector
If multiple I/O requests hit different I/O barrier units, they only need
to compete I/O barrier with other I/Os which hit the same I/O barrier
bucket index with each other. The index of a barrier bucket which a
bio should look for is calculated by sector_to_idx() which is defined
in raid1.h as an inline function,
static inline int sector_to_idx(sector_t sector)
{
return hash_long(sector >> BARRIER_UNIT_SECTOR_BITS,
BARRIER_BUCKETS_NR_BITS);
}
Here sector_nr is the start sector number of a bio.
- Single bio won't go across boundary of a I/O barrier unit
If a request goes across boundary of barrier unit, it will be split. A
bio may be split in raid1_make_request() or raid1_sync_request(), if
sectors returned by align_to_barrier_unit_end() is smaller than
original bio size.
Comparing to single sliding resync window,
- Currently resync I/O grows linearly, therefore regular and resync I/O
will conflict within a single barrier units. So the I/O behavior is
similar to single sliding resync window.
- But a barrier unit bucket is shared by all barrier units with identical
barrier uinit index, the probability of conflict might be higher
than single sliding resync window, in condition that writing I/Os
always hit barrier units which have identical barrier bucket indexs with
the resync I/Os. This is a very rare condition in real I/O work loads,
I cannot imagine how it could happen in practice.
- Therefore we can achieve a good enough low conflict rate with much
simpler barrier algorithm and implementation.
There are two changes should be noticed,
- In raid1d(), I change the code to decrease conf->nr_pending[idx] into
single loop, it looks like this,
spin_lock_irqsave(&conf->device_lock, flags);
conf->nr_queued[idx]--;
spin_unlock_irqrestore(&conf->device_lock, flags);
This change generates more spin lock operations, but in next patch of
this patch set, it will be replaced by a single line code,
atomic_dec(&conf->nr_queueud[idx]);
So we don't need to worry about spin lock cost here.
- Mainline raid1 code split original raid1_make_request() into
raid1_read_request() and raid1_write_request(). If the original bio
goes across an I/O barrier unit size, this bio will be split before
calling raid1_read_request() or raid1_write_request(), this change
the code logic more simple and clear.
- In this patch wait_barrier() is moved from raid1_make_request() to
raid1_write_request(). In raid_read_request(), original wait_barrier()
is replaced by raid1_read_request().
The differnece is wait_read_barrier() only waits if array is frozen,
using different barrier function in different code path makes the code
more clean and easy to read.
Changelog
V4:
- Add alloc_r1bio() to remove redundant r1bio memory allocation code.
- Fix many typos in patch comments.
- Use (PAGE_SHIFT - ilog2(sizeof(int))) to define BARRIER_BUCKETS_NR_BITS.
V3:
- Rebase the patch against latest upstream kernel code.
- Many fixes by review comments from Neil,
- Back to use pointers to replace arraries in struct r1conf
- Remove total_barriers from struct r1conf
- Add more patch comments to explain how/why the values of
BARRIER_UNIT_SECTOR_SIZE and BARRIER_BUCKETS_NR are decided.
- Use get_unqueued_pending() to replace get_all_pendings() and
get_all_queued()
- Increase bucket number from 512 to 1024
- Change code comments format by review from Shaohua.
V2:
- Use bio_split() to split the orignal bio if it goes across barrier unit
bounday, to make the code more simple, by suggestion from Shaohua and
Neil.
- Use hash_long() to replace original linear hash, to avoid a possible
confilict between resync I/O and sequential write I/O, by suggestion from
Shaohua.
- Add conf->total_barriers to record barrier depth, which is used to
control number of parallel sync I/O barriers, by suggestion from Shaohua.
- In V1 patch the bellowed barrier buckets related members in r1conf are
allocated in memory page. To make the code more simple, V2 patch moves
the memory space into struct r1conf, like this,
- int nr_pending;
- int nr_waiting;
- int nr_queued;
- int barrier;
+ int nr_pending[BARRIER_BUCKETS_NR];
+ int nr_waiting[BARRIER_BUCKETS_NR];
+ int nr_queued[BARRIER_BUCKETS_NR];
+ int barrier[BARRIER_BUCKETS_NR];
This change is by the suggestion from Shaohua.
- Remove some inrelavent code comments, by suggestion from Guoqing.
- Add a missing wait_barrier() before jumping to retry_write, in
raid1_make_write_request().
V1:
- Original RFC patch for comments
Signed-off-by: Coly Li <colyli@suse.de>
Cc: Johannes Thumshirn <jthumshirn@suse.de>
Cc: Guoqing Jiang <gqjiang@suse.com>
Reviewed-by: Neil Brown <neilb@suse.de>
Signed-off-by: Shaohua Li <shli@fb.com>
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Although llist provides proper APIs, they are not used. Make them used.
Signed-off-by: Byungchul Park <byungchul.park@lge.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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Firstly bio_clone_mddev() is used in raid normal I/O and isn't
in resync I/O path.
Secondly all the direct access to bvec table in raid happens on
resync I/O except for write behind of raid1, in which we still
use bio_clone() for allocating new bvec table.
So this patch replaces bio_clone() with bio_clone_fast()
in bio_clone_mddev().
Also kill bio_clone_mddev() and call bio_clone_fast() directly, as
suggested by Christoph Hellwig.
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Ming Lei <tom.leiming@gmail.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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mddev is never NULL and neither is ->bio_set, so
remove the check.
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Ming Lei <tom.leiming@gmail.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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Write behind need to replace pages in bio's bvecs, and we have
to clone a fresh bio with new bvec table, so use the introduced
bio_clone_bioset_partial() for it.
For other bio_clone_mddev() cases, we will use fast clone since
they don't need to touch bvec table.
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Ming Lei <tom.leiming@gmail.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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The current behaviour is to fall back to allocate
bio from 'fs_bio_set', that isn't a correct way
because it might cause deadlock.
So this patch simply return failure if mddev->bio_set
can't be created.
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Ming Lei <tom.leiming@gmail.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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md still need bio clone(not the fast version) for behind write,
and it is more efficient to use bio_clone_bioset_partial().
The idea is simple and just copy the bvecs range specified from
parameters.
Reviewed-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Jens Axboe <axboe@fb.com>
Signed-off-by: Ming Lei <tom.leiming@gmail.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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This makes md do the same thing as dm for write same IO failure. Please
see 7eee4ae(dm: disable WRITE SAME if it fails) for details why we need
this.
We did a little bit different than dm. Instead of disabling writesame in
the first IO error, we disable it till next writesame IO coming after
the first IO error. This way we don't need to clone a bio.
Also reported here: https://bugzilla.kernel.org/show_bug.cgi?id=118581
Suggested-by: NeilBrown <neilb@suse.com>
Acked-by: NeilBrown <neilb@suse.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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stripes which are being reclaimed are still accounted into cached
stripes. The reclaim takes time. r5c_do_reclaim isn't aware of the
stripes and does unnecessary stripe reclaim. In practice, I saw one
stripe is reclaimed one time. This will cause bad IO pattern. Fixing
this by excluding the reclaing stripes in the check.
Cc: Song Liu <songliubraving@fb.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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When log space is tight, we try to reclaim stripes from log head. There
are stripes which can't be reclaimed right now if some conditions are
met. We skip such stripes but accidentally count them, which might cause
no stripes are claimed. Fixing this by only counting valid stripes.
Cc: Song Liu <songliubraving@fb.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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I'm starting document of the raid5-cache feature. Please note this is a
kernel doc instead of a mdadm manual, so I don't add the details about
how to use the feature in mdadm side.
Cc: NeilBrown <neilb@suse.com>
Reviewed-by: Song Liu <songliubraving@fb.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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Signed-off-by: Shaohua Li <shli@fb.com>
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Commit: cbd199837750 ("md: Fix unfortunate interaction with evms")
change mddev_put() so that it would not destroy an md device while
->ctime was non-zero.
Unfortunately, we didn't make sure to clear ->ctime when unloading
the module, so it is possible for an md device to remain after
module unload. An attempt to open such a device will trigger
an invalid memory reference in:
get_gendisk -> kobj_lookup -> exact_lock -> get_disk
when tring to access disk->fops, which was in the module that has
been removed.
So ensure we clear ->ctime in md_exit(), and explain how that is useful,
as it isn't immediately obvious when looking at the code.
Fixes: cbd199837750 ("md: Fix unfortunate interaction with evms")
Tested-by: Guoqing Jiang <gqjiang@suse.com>
Signed-off-by: NeilBrown <neilb@suse.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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It is important to be able to flush all stripes in raid5-cache.
Therefore, we need reserve some space on the journal device for
these flushes. If flush operation includes pending writes to the
stripe, we need to reserve (conf->raid_disk + 1) pages per stripe
for the flush out. This reduces the efficiency of journal space.
If we exclude these pending writes from flush operation, we only
need (conf->max_degraded + 1) pages per stripe.
With this patch, when log space is critical (R5C_LOG_CRITICAL=1),
pending writes will be excluded from stripe flush out. Therefore,
we can reduce reserved space for flush out and thus improve journal
device efficiency.
Signed-off-by: Song Liu <songliubraving@fb.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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Chunk aligned read significantly reduces CPU usage of raid456.
However, it is not safe to fully bypass the write back cache.
This patch enables chunk aligned read with write back cache.
For chunk aligned read, we track stripes in write back cache at
a bigger granularity, "big_stripe". Each chunk may contain more
than one stripe (for example, a 256kB chunk contains 64 4kB-page,
so this chunk contain 64 stripes). For chunk_aligned_read, these
stripes are grouped into one big_stripe, so we only need one lookup
for the whole chunk.
For each big_stripe, struct big_stripe_info tracks how many stripes
of this big_stripe are in the write back cache. We count how many
stripes of this big_stripe are in the write back cache. These
counters are tracked in a radix tree (big_stripe_tree).
r5c_tree_index() is used to calculate keys for the radix tree.
chunk_aligned_read() calls r5c_big_stripe_cached() to look up
big_stripe of each chunk in the tree. If this big_stripe is in the
tree, chunk_aligned_read() aborts. This look up is protected by
rcu_read_lock().
It is necessary to remember whether a stripe is counted in
big_stripe_tree. Instead of adding new flag, we reuses existing flags:
STRIPE_R5C_PARTIAL_STRIPE and STRIPE_R5C_FULL_STRIPE. If either of these
two flags are set, the stripe is counted in big_stripe_tree. This
requires moving set_bit(STRIPE_R5C_PARTIAL_STRIPE) to
r5c_try_caching_write(); and moving clear_bit of
STRIPE_R5C_PARTIAL_STRIPE and STRIPE_R5C_FULL_STRIPE to
r5c_finish_stripe_write_out().
Signed-off-by: Song Liu <songliubraving@fb.com>
Reviewed-by: NeilBrown <neilb@suse.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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It will be used in drivers/md/raid5-cache.c
Signed-off-by: Song Liu <songliubraving@fb.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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We made raid5 stripe handling multi-thread before. It works well for
SSD. But for harddisk, the multi-threading creates more disk seek, so
not always improve performance. For several hard disks based raid5,
multi-threading is required as raid5d becames a bottleneck especially
for sequential write.
To overcome the disk seek issue, we only dispatch IO from raid5d if the
array is harddisk based. Other threads can still handle stripes, but
can't dispatch IO.
Idealy, we should control IO dispatching order according to IO position
interrnally. Right now we still depend on block layer, which isn't very
efficient sometimes though.
My setup has 9 harddisks, each disk can do around 180M/s sequential
write. So in theory, the raid5 can do 180 * 8 = 1440M/s sequential
write. The test machine uses an ATOM CPU. I measure sequential write
with large iodepth bandwidth to raid array:
without patch: ~600M/s
without patch and group_thread_cnt=4: 750M/s
with patch and group_thread_cnt=4: 950M/s
with patch, group_thread_cnt=4, skip_copy=1: 1150M/s
We are pretty close to the maximum bandwidth in the large iodepth
iodepth case. The performance gap of small iodepth sequential write
between software raid and theory value is still very big though, because
we don't have an efficient pipeline.
Cc: NeilBrown <neilb@suse.com>
Cc: Song Liu <songliubraving@fb.com>
Signed-off-by: Shaohua Li <shli@fb.com>
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Recently I receive a bug report that on Linux v3.0 based kerenl, hot add
disk to a md linear device causes kernel crash at linear_congested(). From
the crash image analysis, I find in linear_congested(), mddev->raid_disks
contains value N, but conf->disks[] only has N-1 pointers available. Then
a NULL pointer deference crashes the kernel.
There is a race between linear_add() and linear_congested(), RCU stuffs
used in these two functions cannot avoid the race. Since Linuv v4.0
RCU code is replaced by introducing mddev_suspend(). After checking the
upstream code, it seems linear_congested() is not called in
generic_make_request() code patch, so mddev_suspend() cannot provent it
from being called. The possible race still exists.
Here I explain how the race still exists in current code. For a machine
has many CPUs, on one CPU, linear_add() is called to add a hard disk to a
md linear device; at the same time on other CPU, linear_congested() is
called to detect whether this md linear device is congested before issuing
an I/O request onto it.
Now I use a possible code execution time sequence to demo how the possible
race happens,
seq linear_add() linear_congested()
0 conf=mddev->private
1 oldconf=mddev->private
2 mddev->raid_disks++
3 for (i=0; i<mddev->raid_disks;i++)
4 bdev_get_queue(conf->disks[i].rdev->bdev)
5 mddev->private=newconf
In linear_add() mddev->raid_disks is increased in time seq 2, and on
another CPU in linear_congested() the for-loop iterates conf->disks[i] by
the increased mddev->raid_disks in time seq 3,4. But conf with one more
element (which is a pointer to struct dev_info type) to conf->disks[] is
not updated yet, accessing its structure member in time seq 4 will cause a
NULL pointer deference fault.
To fix this race, there are 2 parts of modification in the patch,
1) Add 'int raid_disks' in struct linear_conf, as a copy of
mddev->raid_disks. It is initialized in linear_conf(), always being
consistent with pointers number of 'struct dev_info disks[]'. When
iterating conf->disks[] in linear_congested(), use conf->raid_disks to
replace mddev->raid_disks in the for-loop, then NULL pointer deference
will not happen again.
2) RCU stuffs are back again, and use kfree_rcu() in linear_add() to
free oldconf memory. Because oldconf may be referenced as mddev->private
in linear_congested(), kfree_rcu() makes sure that its memory will not
be released until no one uses it any more.
Also some code comments are added in this patch, to make this modification
to be easier understandable.
This patch can be applied for kernels since v4.0 after commit:
3be260cc18f8 ("md/linear: remove rcu protections in favour of
suspend/resume"). But this bug is reported on Linux v3.0 based kernel, for
people who maintain kernels before Linux v4.0, they need to do some back
back port to this patch.
Changelog:
- V3: add 'int raid_disks' in struct linear_conf, and use kfree_rcu() to
replace rcu_call() in linear_add().
- v2: add RCU stuffs by suggestion from Shaohua and Neil.
- v1: initial effort.
Signed-off-by: Coly Li <colyli@suse.de>
Cc: Shaohua Li <shli@fb.com>
Cc: Neil Brown <neilb@suse.com>
Cc: stable@vger.kernel.org
Signed-off-by: Shaohua Li <shli@fb.com>
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Pull block updates and fixes from Jens Axboe:
- NVMe updates and fixes that missed the first pull request. This
includes bug fixes, and support for autonomous power management.
- Fix from Christoph for missing clear of the request payload, causing
a problem with (at least) the storvsc driver.
- Further fixes for the queue/bdi life time issues from Jan.
- The Kconfig mq scheduler update from me.
- Fixing a use-after-free in dm-rq, spotted by Bart, introduced in this
merge window.
- Three fixes for nbd from Josef.
- Bug fix from Omar, fixing a bug in sas transport code that oopses
when bsg ioctls were used. From Omar.
- Improvements to the queue restart and tag wait from from Omar.
- Set of fixes for the sed/opal code from Scott.
- Three trivial patches to cciss from Tobin
* 'for-linus' of git://git.kernel.dk/linux-block: (41 commits)
dm-rq: don't dereference request payload after ending request
blk-mq-sched: separate mark hctx and queue restart operations
blk-mq: use sbq wait queues instead of restart for driver tags
block/sed-opal: Propagate original error message to userland.
nvme/pci: re-check security protocol support after reset
block/sed-opal: Introduce free_opal_dev to free the structure and clean up state
nvme: detect NVMe controller in recent MacBooks
nvme-rdma: add support for host_traddr
nvmet-rdma: Fix error handling
nvmet-rdma: use nvme cm status helper
nvme-rdma: move nvme cm status helper to .h file
nvme-fc: don't bother to validate ioccsz and iorcsz
nvme/pci: No special case for queue busy on IO
nvme/core: Fix race kicking freed request_queue
nvme/pci: Disable on removal when disconnected
nvme: Enable autonomous power state transitions
nvme: Add a quirk mechanism that uses identify_ctrl
nvme: make nvmf_register_transport require a create_ctrl callback
nvme: Use CNS as 8-bit field and avoid endianness conversion
nvme: add semicolon in nvme_command setting
...
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Bart reported a case where dm would crash with use-after-free
poison. This is due to dm_softirq_done() accessing memory
associated with a request after calling end_request on it.
This is most visible on !blk-mq, since we free the memory
immediately for that case.
Reported-by: Bart Van Assche <bart.vanassche@sandisk.com>
Suggested-by: Linus Torvalds <torvalds@linux-foundation.org>
Fixes: eb8db831be80 ("dm: always defer request allocation to the owner of the request_queue")
Signed-off-by: Jens Axboe <axboe@fb.com>
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In blk_mq_sched_dispatch_requests(), we call blk_mq_sched_mark_restart()
after we dispatch requests left over on our hardware queue dispatch
list. This is so we'll go back and dispatch requests from the scheduler.
In this case, it's only necessary to restart the hardware queue that we
are running; there's no reason to run other hardware queues just because
we are using shared tags.
So, split out blk_mq_sched_mark_restart() into two operations, one for
just the hardware queue and one for the whole request queue. The core
code only needs the hctx variant, but I/O schedulers will want to use
both.
This also requires adjusting blk_mq_sched_restart_queues() to always
check the queue restart flag, not just when using shared tags.
Signed-off-by: Omar Sandoval <osandov@fb.com>
Signed-off-by: Jens Axboe <axboe@fb.com>
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Commit 50e1dab86aa2 ("blk-mq-sched: fix starvation for multiple hardware
queues and shared tags") fixed one starvation issue for shared tags.
However, we can still get into a situation where we fail to allocate a
tag because all tags are allocated but we don't have any pending
requests on any hardware queue.
One solution for this would be to restart all queues that share a tag
map, but that really sucks. Ideally, we could just block and wait for a
tag, but that isn't always possible from blk_mq_dispatch_rq_list().
However, we can still use the struct sbitmap_queue wait queues with a
custom callback instead of blocking. This has a few benefits:
1. It avoids iterating over all hardware queues when completing an I/O,
which the current restart code has to do.
2. It benefits from the existing rolling wakeup code.
3. It avoids punting to another thread just to have it block.
Signed-off-by: Omar Sandoval <osandov@fb.com>
Signed-off-by: Jens Axboe <axboe@fb.com>
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During an error on a comannd, ex: user provides wrong pw to unlock
range, we will gracefully terminate the opal session. We want to
propagate the original error to userland instead of the result of
the session termination, which is almost always a success.
Signed-off-by: Scott Bauer <scott.bauer@intel.com>
Signed-off-by: Jens Axboe <axboe@fb.com>
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A device may change capabilities after each reset, e.g. due to a firmware
upgrade. We should thus check for Security Send/Receive and OPAL support
after each reset.
Based on patches from Christoph and Keith.
Signed-off-by: Scott Bauer <scott.bauer@intel.com>
Signed-off-by: Jens Axboe <axboe@fb.com>
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Before we free the opal structure we need to clean up any saved
locking ranges that the user had told us to unlock from a suspend.
Signed-off-by: Scott Bauer <scott.bauer@intel.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Jens Axboe <axboe@fb.com>
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