The reverted changes were intended to force a choice of whether any
newly-added BufferGetPage() calls needed to be accompanied by a
test of the snapshot age, to support the "snapshot too old"
feature. Such an accompanying test is needed in about 7% of the
cases, where the page is being used as part of a scan rather than
positioning for other purposes (such as DML or vacuuming). The
additional effort required for back-patching, and the doubt whether
the intended benefit would really be there, have indicated it is
best just to rely on developers to do the right thing based on
comments and existing usage, as we do with many other conventions.
This change should have little or no effect on generated executable
code.
Motivated by the back-patching pain of Tom Lane and Robert Haas
Coverity complained that oldPartitionLock was possibly dereferenced after
having been set to NULL. That actually can't happen, because we'd only use
it if (oldFlags & BM_TAG_VALID) is true. But nonetheless Coverity is
justified in complaining, because at line 1275 we actually overwrite
oldFlags, and then still expect its BM_TAG_VALID bit to be a safe guide to
whether to release the oldPartitionLock. Thus, the code would be incorrect
if someone else had changed the buffer's BM_TAG_VALID flag meanwhile.
That should not happen, since we hold pin on the buffer throughout this
sequence, but it's starting to look like a rather shaky chain of logic.
And there's no need for such assumptions, because we can simply replace
the (oldFlags & BM_TAG_VALID) tests with (oldPartitionLock != NULL),
which has identical results and makes it plain to all comers that we don't
dereference a null pointer. A small side benefit is that the range of
liveness of oldFlags is greatly reduced, possibly allowing the compiler
to save a register.
This is just cleanup, not an actual bug fix, so there seems no need
for a back-patch.
We've had repeated troubles over the years with failures to initialize
spinlocks correctly; see 6b93fcd14 for a recent example. Most of the time,
on most platforms, such oversights can escape notice because all-zeroes is
the expected initial content of an slock_t variable. The only platform
we have where the initialized state of an slock_t isn't zeroes is HPPA,
and that's practically gone in the wild. To make it easier to catch such
errors without needing one of those, adjust the --disable-spinlocks code
so that zero is not a valid value for an slock_t for it.
In passing, remove a bunch of unnecessary #include's from spin.c;
commit daa7527afc removed all the intermodule coupling that
made them necessary.
pg_xlogdump includes bufmgr.h. With a compiler that emits code for
static inline functions even when they're unreferenced, that leads
to unresolved external references in the new static-inline version
of BufferGetPage(). So hide it with #ifndef FRONTEND, as we've done
for similar issues elsewhere. Per buildfarm member pademelon.
My previous attempt at doing so, in 80abbeba23, was not sufficient. While that
fixed the problem for bufmgr.c and lwlock.c , s_lock.c still has non-constant
expressions in the struct initializer, because the file/line/function
information comes from the caller of s_lock().
Give up on using a macro, and use a static inline instead.
Discussion: 4369.1460435533@sss.pgh.pa.us
The current definition of init_spin_delay (introduced recently in
48354581a) wasn't C89 compliant. It's not legal to refer to refer to
non-constant expressions, and the ptr argument was one. This, as
reported by Tom, lead to a failure on buildfarm animal pademelon.
The pointer, especially on system systems with ASLR, isn't super helpful
anyway, though. So instead of making init_spin_delay into an inline
function, make s_lock_stuck() report the function name in addition to
file:line and change init_spin_delay() accordingly. While not a direct
replacement, the function name is likely more useful anyway (line
numbers are often hard to interpret in third party reports).
This also fixes what file/line number is reported for waits via
s_lock().
As PG_FUNCNAME_MACRO is now used outside of elog.h, move it to c.h.
Reported-By: Tom Lane
Discussion: 4369.1460435533@sss.pgh.pa.us
The recent patch to make Pin/UnpinBuffer lockfree in the hot
path (48354581a), accidentally used pg_atomic_fetch_or_u32() in
MarkLocalBufferDirty(). Other code operating on local buffers was
careful to only use pg_atomic_read/write_u32 which just read/write from
memory; to avoid unnecessary overhead.
On its own that'd just make MarkLocalBufferDirty() slightly less
efficient, but in addition InitLocalBuffers() doesn't call
pg_atomic_init_u32() - thus the spinlock fallback for the atomic
operations isn't initialized. That in turn caused, as reported by Tom,
buildfarm animal gaur to fail. As those errors are actually useful
against this type of error, continue to omit - intentionally this time -
initialization of the atomic variable.
In addition, add an explicit note about only using pg_atomic_read/write
on local buffers's state to BufferDesc's description.
Reported-By: Tom Lane
Discussion: 1881.1460431476@sss.pgh.pa.us
It's silly to define these counts as narrower than they might someday
need to be. Also, I believe that the BLCKSZ * nflush calculation in
mdwriteback was capable of overflowing an int.
Commit 428b1d6b29 introduced the use of
msync() for flushing dirty data from the kernel's file buffers. Several
portability issues were overlooked, though:
* Not all implementations of mmap() think that nbytes == 0 means "map
the whole file". To fix, use lseek() to find out the true length.
Fix callers of pg_flush_data to be aware that nbytes == 0 may result
in trashing the file's seek position.
* Not all implementations of mmap() will accept partial-page mmap
requests. To fix, round down the length request to whatever sysconf()
says the page size is. (I think this is OK from a portability standpoint,
because sysconf() is required by SUS v2, and we aren't trying to compile
this part on Windows anyway. Buildfarm should let us know if not.)
* On 32-bit machines, the file size might exceed the available free
address space, or even exceed what will fit in size_t. Check for
the latter explicitly to avoid passing a false request size to mmap().
If mmap fails, silently fall through to the next implementation method,
rather than bleating to the postmaster log and giving up.
* mmap'ing directories fails on some platforms, and even if it works,
msync'ing the directory is quite unlikely to help, as for that matter are
the other flush implementations. In pre_sync_fname(), just skip flush
attempts on directories.
In passing, copy-edit the comments a bit.
Stas Kelvich and myself
On a big NUMA machine with 1000 connections in saturation load
there was a performance regression due to spinlock contention, for
acquiring values which were never used. Just fill with dummy
values if we're not going to use them.
This patch has not been benchmarked yet on a big NUMA machine, but
it seems like a good idea on general principle, and it seemed to
prevent an apparent 2.2% regression on a single-socket i7 box
running 200 connections at saturation load.
I was initially concerned that the some of the hundreds of
references to BufferGetPage() where the literal
BGP_NO_SNAPSHOT_TEST were passed might not optimize as well as a
macro, leading to some hard-to-find performance regressions in
corner cases. Inspection of disassembled code has shown identical
code at all inspected locations, and the size difference doesn't
amount to even one byte per such call. So make it readable.
Per gripes from Álvaro Herrera and Tom Lane
Previously we used a spinlock, in adition to the atomically manipulated
->state field, to protect the wait queue. But it's pretty simple to
instead perform the locking using a flag in state.
Due to 6150a1b0 BufferDescs, on platforms (like PPC) with > 1 byte
spinlocks, increased their size above 64byte. As 64 bytes are the size
we pad allocated BufferDescs to, this can increase false sharing;
causing performance problems in turn. Together with the previous commit
this reduces the size to <= 64 bytes on all common platforms.
Author: Andres Freund
Discussion: CAA4eK1+ZeB8PMwwktf+3bRS0Pt4Ux6Rs6Aom0uip8c6shJWmyg@mail.gmail.com20160327121858.zrmrjegmji2ymnvr@alap3.anarazel.de
Pinning/Unpinning a buffer is a very frequent operation; especially in
read-mostly cache resident workloads. Benchmarking shows that in various
scenarios the spinlock protecting a buffer header's state becomes a
significant bottleneck. The problem can be reproduced with pgbench -S on
larger machines, but can be considerably worse for queries which touch
the same buffers over and over at a high frequency (e.g. nested loops
over a small inner table).
To allow atomic operations to be used, cram BufferDesc's flags,
usage_count, buf_hdr_lock, refcount into a single 32bit atomic variable;
that allows to manipulate them together using 32bit compare-and-swap
operations. This requires reducing MAX_BACKENDS to 2^18-1 (which could
be lifted by using a 64bit field, but it's not a realistic configuration
atm).
As not all operations can easily implemented in a lockfree manner,
implement the previous buf_hdr_lock via a flag bit in the atomic
variable. That way we can continue to lock the header in places where
it's needed, but can get away without acquiring it in the more frequent
hot-paths. There's some additional operations which can be done without
the lock, but aren't in this patch; but the most important places are
covered.
As bufmgr.c now essentially re-implements spinlocks, abstract the delay
logic from s_lock.c into something more generic. It now has already two
users, and more are coming up; there's a follupw patch for lwlock.c at
least.
This patch is based on a proof-of-concept written by me, which Alexander
Korotkov made into a fully working patch; the committed version is again
revised by me. Benchmarking and testing has, amongst others, been
provided by Dilip Kumar, Alexander Korotkov, Robert Haas.
On a large x86 system improvements for readonly pgbench, with a high
client count, of a factor of 8 have been observed.
Author: Alexander Korotkov and Andres Freund
Discussion: 2400449.GjM57CE0Yg@dinodell
This feature is controlled by a new old_snapshot_threshold GUC. A
value of -1 disables the feature, and that is the default. The
value of 0 is just intended for testing. Above that it is the
number of minutes a snapshot can reach before pruning and vacuum
are allowed to remove dead tuples which the snapshot would
otherwise protect. The xmin associated with a transaction ID does
still protect dead tuples. A connection which is using an "old"
snapshot does not get an error unless it accesses a page modified
recently enough that it might not be able to produce accurate
results.
This is similar to the Oracle feature, and we use the same SQLSTATE
and error message for compatibility.
This patch is a no-op patch which is intended to reduce the chances
of failures of omission once the functional part of the "snapshot
too old" patch goes in. It adds parameters for snapshot, relation,
and an enum to specify whether the snapshot age check needs to be
done for the page at this point. This initial patch passes NULL
for the first two new parameters and BGP_NO_SNAPSHOT_TEST for the
third. The follow-on patch will change the places where the test
needs to be made.
Contention on the relation extension lock can become quite fierce when
multiple processes are inserting data into the same relation at the same
time at a high rate. Experimentation shows the extending the relation
multiple blocks at a time improves scalability.
Dilip Kumar, reviewed by Petr Jelinek, Amit Kapila, and me.
Experimentation shows this only costs about 6kB, which seems well
worth it given the major performance effects that can be caused
by insufficient alignment, especially on larger systems.
Discussion: 14166.1458924422@sss.pgh.pa.us
If archive_timeout > 0 we should avoid logging XLOG_RUNNING_XACTS if idle.
Bug 13685 reported by Laurence Rowe, investigated in detail by Michael Paquier,
though this is not his proposed fix.
20151016203031.3019.72930@wrigleys.postgresql.org
Simple non-invasive patch to allow later backpatch to 9.4 and 9.5
Commit ac1d794 ("Make idle backends exit if the postmaster dies.")
introduced a regression on, at least, large linux systems. Constantly
adding the same postmaster_alive_fds to the OSs internal datastructures
for implementing poll/select can cause significant contention; leading
to a performance regression of nearly 3x in one example.
This can be avoided by using e.g. linux' epoll, which avoids having to
add/remove file descriptors to the wait datastructures at a high rate.
Unfortunately the current latch interface makes it hard to allocate any
persistent per-backend resources.
Replace, with a backward compatibility layer, WaitLatchOrSocket with a
new WaitEventSet API. Users can allocate such a Set across multiple
calls, and add more than one file-descriptor to wait on. The latter has
been added because there's upcoming postgres features where that will be
helpful.
In addition to the previously existing poll(2), select(2),
WaitForMultipleObjects() implementations also provide an epoll_wait(2)
based implementation to address the aforementioned performance
problem. Epoll is only available on linux, but that is the most likely
OS for machines large enough (four sockets) to reproduce the problem.
To actually address the aforementioned regression, create and use a
long-lived WaitEventSet for FE/BE communication. There are additional
places that would benefit from a long-lived set, but that's a task for
another day.
Thanks to Amit Kapila, who helped make the windows code I blindly wrote
actually work.
Reported-By: Dmitry Vasilyev Discussion:
CAB-SwXZh44_2ybvS5Z67p_CDz=XFn4hNAD=CnMEF+QqkXwFrGg@mail.gmail.com20160114143931.GG10941@awork2.anarazel.de
Previously latches for windows and unix had been implemented in
different files. A later patch introduce an expanded wait
infrastructure, keeping the implementation separate would introduce too
much duplication.
This basically just moves the functions, without too much change. The
reason to keep this separate is that it allows blame to continue working
a little less badly; and to make review a tiny bit easier.
Discussion: 20160114143931.GG10941@awork2.anarazel.de
Buildfarm members gaur and pademelon are old enough not to know about
MAP_FAILED; which is used in 428b1d6. Include portability/mem.h to fix;
as already done in a bunch of other places.
The old code is wrong, because it returns a pointer to an automatic
variable. And it's also more clever than we really need to be
considering that the case it's worrying about should never happen.
Up to now checkpoints were written in the order they're in the
BufferDescriptors. That's nearly random in a lot of cases, which
performs badly on rotating media, but even on SSDs it causes slowdowns.
To avoid that, sort checkpoints before writing them out. We currently
sort by tablespace, relfilenode, fork and block number.
One of the major reasons that previously wasn't done, was fear of
imbalance between tablespaces. To address that balance writes between
tablespaces.
The other prime concern was that the relatively large allocation to sort
the buffers in might fail, preventing checkpoints from happening. Thus
pre-allocate the required memory in shared memory, at server startup.
This particularly makes it more efficient to have checkpoint flushing
enabled, because that'll often result in a lot of writes that can be
coalesced into one flush.
Discussion: alpine.DEB.2.10.1506011320000.28433@sto
Author: Fabien Coelho and Andres Freund
Currently writes to the main data files of postgres all go through the
OS page cache. This means that some operating systems can end up
collecting a large number of dirty buffers in their respective page
caches. When these dirty buffers are flushed to storage rapidly, be it
because of fsync(), timeouts, or dirty ratios, latency for other reads
and writes can increase massively. This is the primary reason for
regular massive stalls observed in real world scenarios and artificial
benchmarks; on rotating disks stalls on the order of hundreds of seconds
have been observed.
On linux it is possible to control this by reducing the global dirty
limits significantly, reducing the above problem. But global
configuration is rather problematic because it'll affect other
applications; also PostgreSQL itself doesn't always generally want this
behavior, e.g. for temporary files it's undesirable.
Several operating systems allow some control over the kernel page
cache. Linux has sync_file_range(2), several posix systems have msync(2)
and posix_fadvise(2). sync_file_range(2) is preferable because it
requires no special setup, whereas msync() requires the to-be-flushed
range to be mmap'ed. For the purpose of flushing dirty data
posix_fadvise(2) is the worst alternative, as flushing dirty data is
just a side-effect of POSIX_FADV_DONTNEED, which also removes the pages
from the page cache. Thus the feature is enabled by default only on
linux, but can be enabled on all systems that have any of the above
APIs.
While desirable and likely possible this patch does not contain an
implementation for windows.
With the infrastructure added, writes made via checkpointer, bgwriter
and normal user backends can be flushed after a configurable number of
writes. Each of these sources of writes controlled by a separate GUC,
checkpointer_flush_after, bgwriter_flush_after and backend_flush_after
respectively; they're separate because the number of flushes that are
good are separate, and because the performance considerations of
controlled flushing for each of these are different.
A later patch will add checkpoint sorting - after that flushes from the
ckeckpoint will almost always be desirable. Bgwriter flushes are most of
the time going to be random, which are slow on lots of storage hardware.
Flushing in backends works well if the storage and bgwriter can keep up,
but if not it can have negative consequences. This patch is likely to
have negative performance consequences without checkpoint sorting, but
unfortunately so has sorting without flush control.
Discussion: alpine.DEB.2.10.1506011320000.28433@sto
Author: Fabien Coelho and Andres Freund
Earlier version committed in 9.0 caused spurious waits in some cases.
New infrastructure for lock waits in 9.3 used to correct and improve this.
Jeff Janes based upon a proposal by Simon Riggs, who also reviewed
Additional review comments from Amit Kapila
When a process is waiting for a heavyweight lock, we will now indicate
the type of heavyweight lock for which it is waiting. Also, you can
now see when a process is waiting for a lightweight lock - in which
case we will indicate the individual lock name or the tranche, as
appropriate - or for a buffer pin.
Amit Kapila, Ildus Kurbangaliev, reviewed by me. Lots of helpful
discussion and suggestions by many others, including Alexander
Korotkov, Vladimir Borodin, and many others.
Renaming a file using rename(2) is not guaranteed to be durable in face
of crashes; especially on filesystems like xfs and ext4 when mounted
with data=writeback. To be certain that a rename() atomically replaces
the previous file contents in the face of crashes and different
filesystems, one has to fsync the old filename, rename the file, fsync
the new filename, fsync the containing directory. This sequence is not
generally adhered to currently; which exposes us to data loss risks. To
avoid having to repeat this arduous sequence, introduce
durable_rename(), which wraps all that.
Also add durable_link_or_rename(). Several places use link() (with a
fallback to rename()) to rename a file, trying to avoid replacing the
target file out of paranoia. Some of those rename sequences need to be
durable as well. There seems little reason extend several copies of the
same logic, so centralize the link() callers.
This commit does not yet make use of the new functions; they're used in
a followup commit.
Author: Michael Paquier, Andres Freund
Discussion: 56583BDD.9060302@2ndquadrant.com
Backpatch: All supported branches
For example, if you want to perform an ioctl() on a file descriptor
opened through the fd.c routines, there's no way to do that without
being able to get at the underlying fd.
KaiGai Kohei
The new bit indicates whether every tuple on the page is already frozen.
It is cleared only when the all-visible bit is cleared, and it can be
set only when we vacuum a page and find that every tuple on that page is
both visible to every transaction and in no need of any future
vacuuming.
A future commit will use this new bit to optimize away full-table scans
that would otherwise be triggered by XID wraparound considerations. A
page which is merely all-visible must still be scanned in that case, but
a page which is all-frozen need not be. This commit does not attempt
that optimization, although that optimization is the goal here. It
seems better to get the basic infrastructure in place first.
Per discussion, it's very desirable for pg_upgrade to automatically
migrate existing VM forks from the old format to the new format. That,
too, will be handled in a follow-on patch.
Masahiko Sawada, reviewed by Kyotaro Horiguchi, Fujii Masao, Amit
Kapila, Simon Riggs, Andres Freund, and others, and substantially
revised by me.
This patch introduces "pg_blocking_pids(int) returns int[]", which returns
the PIDs of any sessions that are blocking the session with the given PID.
Historically people have obtained such information using a self-join on
the pg_locks view, but it's unreasonably tedious to do it that way with any
modicum of correctness, and the addition of parallel queries has pretty
much broken that approach altogether. (Given some more columns in the view
than there are today, you could imagine handling parallel-query cases with
a 4-way join; but ugh.)
The new function has the following behaviors that are painful or impossible
to get right via pg_locks:
1. Correctly understands which lock modes block which other ones.
2. In soft-block situations (two processes both waiting for conflicting lock
modes), only the one that's in front in the wait queue is reported to
block the other.
3. In parallel-query cases, reports all sessions blocking any member of
the given PID's lock group, and reports a session by naming its leader
process's PID, which will be the pg_backend_pid() value visible to
clients.
The motivation for doing this right now is mostly to fix the isolation
tests. Commit 38f8bdcac4 lobotomized
isolationtester's is-it-waiting query by removing its ability to recognize
nonconflicting lock modes, as a crude workaround for the inability to
handle soft-block situations properly. But even without the lock mode
tests, the old query was excessively slow, particularly in
CLOBBER_CACHE_ALWAYS builds; some of our buildfarm animals fail the new
deadlock-hard test because the deadlock timeout elapses before they can
probe the waiting status of all eight sessions. Replacing the pg_locks
self-join with use of pg_blocking_pids() is not only much more correct, but
a lot faster: I measure it at about 9X faster in a typical dev build with
Asserts, and 3X faster in CLOBBER_CACHE_ALWAYS builds. That should provide
enough headroom for the slower CLOBBER_CACHE_ALWAYS animals to pass the
test, without having to lengthen deadlock_timeout yet more and thus slow
down the test for everyone else.
We don't really need this field, because it's either zero or redundant with
PGPROC.pid. The use of zero to mark "not a group leader" is not necessary
since we can just as well test whether lockGroupLeader is NULL. This does
not save very much, either as to code or data, but the simplification seems
worthwhile anyway.
In 4b4b680c3 I accidentally used sizeof(PrivateRefCountArray) instead of
sizeof(PrivateRefCountEntry) when creating the refcount overflow
hashtable. As the former is bigger than the latter, this luckily only
resulted in a slightly increased memory usage when many buffers are
pinned in a backend.
Reported-By: Takashi Horikawa
Discussion: 73FA3881462C614096F815F75628AFCD035A48C3@BPXM01GP.gisp.nec.co.jp
Backpatch: 9.5, where thew new ref count infrastructure was introduced
Reflow text in lock manager README so that it fits within 80 columns.
Correct some mistakes. Expand the README to explain not only why group
locking exists but also the data structures that support it. Improve
comments related to group locking several files. Change the name of a
macro argument for improved clarity.
Most of these problems were reported by Tom Lane, but I found a few
of them myself.
Robert Haas and Tom Lane
The API spec for this function was changed completely (and for the better)
by commit 3cba8999b3, but it didn't bother
with anything as mundane as updating the comments.
This finishes the work - spread across many commits over the last
several months - of putting each type of lock other than the named
individual locks into a separate tranche.
Amit Kapila
Commit 0e141c0fbb introduced a new
facility to reduce ProcArrayLock contention by clearing several XIDs
from the ProcArray under a single lock acquisition. The names
initially chosen were deemed not to be very good choices, so commit
4aec49899e renamed them. But now it
seems like we still didn't get it right. A pending patch wants to
add similar infrastructure for batching CLOG updates, so the names
need to be clear enough to allow a new set of structure members with
a related purpose.
Amit Kapila
As of commit c1772ad922, there's no
longer any way of requesting additional LWLocks in the main tranche,
so we don't need NumLWLocks() or LWLockAssign() any more. Also,
some of the allocation counters that we had previously aren't needed
any more either.
Amit Kapila
