pg_filedump and other external utility programs are likely to want to be
able to check Postgres page checksums. To avoid messy duplication of code,
move the checksumming functionality into an exported header file, much as
we did awhile back for the CRC code.
In passing, get rid of an unportable assumption that a static char[] array
will be word-aligned, and do some other minor code beautification.
Since commit f21bb9cfb5, this function
ignores the caller-provided length and loops until it finds a
terminator, which GetVirtualXIDsDelayingChkpt() never adds. Restore the
previous loop control logic. In passing, revert the addition of an
unused variable by the same commit, presumably a debugging relic.
AllocateFile(), AllocateDir(), and some sister routines share a small array
for remembering requests, so that the files can be closed on transaction
failure. Previously that array had a fixed size, MAX_ALLOCATED_DESCS (32).
While historically that had seemed sufficient, Steve Toutant pointed out
that this meant you couldn't scan more than 32 file_fdw foreign tables in
one query, because file_fdw depends on the COPY code which uses
AllocateFile(). There are probably other cases, or will be in the future,
where this nonconfigurable limit impedes users.
We can't completely remove any such limit, at least not without a lot of
work, since each such request requires a kernel file descriptor and most
platforms limit the number we can have. (In principle we could
"virtualize" these descriptors, as fd.c already does for the main VFD pool,
but not without an additional layer of overhead and a lot of notational
impact on the calling code.) But we can at least let the array size be
configurable. Hence, change the code to allow up to max_safe_fds/2
allocated file requests. On modern platforms this should allow several
hundred concurrent file_fdw scans, or more if one increases the value of
max_files_per_process. To go much further than that, we'd need to do some
more work on the data structure, since the current code for closing
requests has potentially O(N^2) runtime; but it should still be all right
for request counts in this range.
Back-patch to 9.1 where contrib/file_fdw was introduced.
The array allocated by GetRunningTransactionLocks() needs to be pfree'd
when we're done with it. Otherwise we leak some memory during each
checkpoint, if wal_level = hot_standby. This manifests as memory bloat
in the checkpointer process, or in bgwriter in versions before we made
the checkpointer separate.
Reported and fixed by Naoya Anzai. Back-patch to 9.0 where the issue
was introduced.
In passing, improve comments for GetRunningTransactionLocks(), and add
an Assert that we didn't overrun the palloc'd array.
PathNameOpenFile failed to ensure that the correct value of errno was
returned to its caller after a failure (because it incorrectly supposed
that free() can never change errno). In some cases this would result
in a user-visible failure because an expected ENOENT errno was replaced
with something else. Bogus EINVAL failures have been observed on OS X,
for example.
There were also a couple of places that could mangle an important value
of errno if FDDEBUG was defined. While the usefulness of that debug
support is highly debatable, we might as well make it safe to use,
so add errno save/restore logic to the DO_DB macro.
Per bug #8167 from Nelson Minar, diagnosed by RhodiumToad.
Back-patch to all supported branches.
Isolate checksum calculation to its own module, so that bufpage
knows little if anything about the details of the calculation.
This implementation is a modified FNV-1a hash checksum, details
of which are given in the new checksum.c header comments.
Basic implementation only, so we fix the output value.
Later related commits will add version numbers to pg_control,
compiler optimization flags and memory barriers.
Ants Aasma, reviewed by Jeff Davis and Simon Riggs
We copy the buffer before inserting an XLOG_HINT to avoid WAL CRC errors
caused by concurrent hint writes to buffer while share locked. To make this work
we refactor RestoreBackupBlock() to allow an XLOG_HINT to avoid the normal
path for backup blocks, which assumes the underlying buffer is exclusive locked.
Resulting code completely changes layout of XLOG_HINT WAL records, but
this isn't even beta code, so this is a low impact change.
In passing, avoid taking WALInsertLock for full page writes on checksummed
hints, remove related cruft from XLogInsert() and improve xlog_desc record for
XLOG_HINT.
Andres Freund
Bug report by Fujii Masao, testing by Jeff Janes and Jaime Casanova,
review by Jeff Davis and Simon Riggs. Applied with changes from review
and some comment editing.
The old formula didn't take into account that each WAL sender process needs
a spinlock. We had also already exceeded the fixed number of spinlocks
reserved for misc purposes (10). Bump that to 30.
Backpatch to 9.0, where WAL senders were introduced. If I counted correctly,
9.0 had exactly 10 predefined spinlocks, and 9.1 exceeded that, but bump the
limit in 9.0 too because 10 is uncomfortably close to the edge.
The modern incarnation of md.c is by no means specific to magnetic disk
technology, but every so often we hear from someone who's misled by the
label. Try to clarify that it will work for anything that supports
standard filesystem operations. Per suggestion from Andrew Dunstan.
Checksums are set immediately prior to flush out of shared buffers
and checked when pages are read in again. Hint bit setting will
require full page write when block is dirtied, which causes various
infrastructure changes. Extensive comments, docs and README.
WARNING message thrown if checksum fails on non-all zeroes page;
ERROR thrown but can be disabled with ignore_checksum_failure = on.
Feature enabled by an initdb option, since transition from option off
to option on is long and complex and has not yet been implemented.
Default is not to use checksums.
Checksum used is WAL CRC-32 truncated to 16-bits.
Simon Riggs, Jeff Davis, Greg Smith
Wide input and assistance from many community members. Thank you.
This GUC allows limiting the time spent waiting to acquire any one
heavyweight lock.
In support of this, improve the recently-added timeout infrastructure
to permit efficiently enabling or disabling multiple timeouts at once.
That reduces the performance hit from turning on lock_timeout, though
it's still not zero.
Zoltán Böszörményi, reviewed by Tom Lane,
Stephen Frost, and Hari Babu
Formerly we just Assert'ed that each refcount was zero, which was quick
and easy but failed to provide a good overview of what was wrong.
Change the code so that we'll call PrintBufferLeakWarning() for each
buffer with a nonzero refcount, and then Assert at the end of the loop.
This costs nothing in runtime and might ease diagnosis of some bugs.
Greg Smith, reviewed by Satoshi Nagayasu, further tweaked by me
A materialized view has a rule just like a view and a heap and
other physical properties like a table. The rule is only used to
populate the table, references in queries refer to the
materialized data.
This is a minimal implementation, but should still be useful in
many cases. Currently data is only populated "on demand" by the
CREATE MATERIALIZED VIEW and REFRESH MATERIALIZED VIEW statements.
It is expected that future releases will add incremental updates
with various timings, and that a more refined concept of defining
what is "fresh" data will be developed. At some point it may even
be possible to have queries use a materialized in place of
references to underlying tables, but that requires the other
above-mentioned features to be working first.
Much of the documentation work by Robert Haas.
Review by Noah Misch, Thom Brown, Robert Haas, Marko Tiikkaja
Security review by KaiGai Kohei, with a decision on how best to
implement sepgsql still pending.
This includes backend "COPY TO/FROM PROGRAM '...'" syntax, and corresponding
psql \copy syntax. Like with reading/writing files, the backend version is
superuser-only, and in the psql version, the program is run in the client.
In the passing, the psql \copy STDIN/STDOUT syntax is subtly changed: if you
the stdin/stdout is quoted, it's now interpreted as a filename. For example,
"\copy foo from 'stdin'" now reads from a file called 'stdin', not from
standard input. Before this, there was no way to specify a filename called
stdin, stdout, pstdin or pstdout.
This creates a new function in pgport, wait_result_to_str(), which can
be used to convert the exit status of a process, as returned by wait(3),
to a human-readable string.
Etsuro Fujita, reviewed by Amit Kapila.
This enables non-backend code, such as pg_xlogdump, to use it easily.
The previous location, in src/backend/catalog/catalog.c, made that
essentially impossible because that file depends on many backend-only
facilities; so this needs to live separately.
libpgcommon is a new static library to allow sharing code among the
various frontend programs and backend; this lets us eliminate duplicate
implementations of common routines. We avoid libpgport, because that's
intended as a place for porting issues; per discussion, it seems better
to keep them separate.
The first use case, and the only implemented by this patch, is pg_malloc
and friends, which many frontend programs were already using.
At the same time, we can use this to provide palloc emulation functions
for the frontend; this way, some palloc-using files in the backend can
also be used by the frontend cleanly. To do this, we change palloc() in
the backend to be a function instead of a macro on top of
MemoryContextAlloc(). This was previously believed to cause loss of
performance, but this implementation has been tweaked by Tom and Andres
so that on modern compilers it provides a slight improvement over the
previous one.
This lets us clean up some places that were already with
localized hacks.
Most of the pg_malloc/palloc changes in this patch were authored by
Andres Freund. Zoltán Böszörményi also independently provided a form of
that. libpgcommon infrastructure was authored by Álvaro.
The reason this wasn't supported before was that GiST indexes need an
increasing sequence to detect concurrent page-splits. In a regular WAL-
logged GiST index, the LSN of the page-split record is used for that
purpose, and in a temporary index, we can get away with a backend-local
counter. Neither of those methods works for an unlogged relation.
To provide such an increasing sequence of numbers, create a "fake LSN"
counter that is saved and restored across shutdowns. On recovery, unlogged
relations are blown away, so the counter doesn't need to survive that
either.
Jeevan Chalke, based on discussions with Robert Haas, Tom Lane and me.
This patch introduces two additional lock modes for tuples: "SELECT FOR
KEY SHARE" and "SELECT FOR NO KEY UPDATE". These don't block each
other, in contrast with already existing "SELECT FOR SHARE" and "SELECT
FOR UPDATE". UPDATE commands that do not modify the values stored in
the columns that are part of the key of the tuple now grab a SELECT FOR
NO KEY UPDATE lock on the tuple, allowing them to proceed concurrently
with tuple locks of the FOR KEY SHARE variety.
Foreign key triggers now use FOR KEY SHARE instead of FOR SHARE; this
means the concurrency improvement applies to them, which is the whole
point of this patch.
The added tuple lock semantics require some rejiggering of the multixact
module, so that the locking level that each transaction is holding can
be stored alongside its Xid. Also, multixacts now need to persist
across server restarts and crashes, because they can now represent not
only tuple locks, but also tuple updates. This means we need more
careful tracking of lifetime of pg_multixact SLRU files; since they now
persist longer, we require more infrastructure to figure out when they
can be removed. pg_upgrade also needs to be careful to copy
pg_multixact files over from the old server to the new, or at least part
of multixact.c state, depending on the versions of the old and new
servers.
Tuple time qualification rules (HeapTupleSatisfies routines) need to be
careful not to consider tuples with the "is multi" infomask bit set as
being only locked; they might need to look up MultiXact values (i.e.
possibly do pg_multixact I/O) to find out the Xid that updated a tuple,
whereas they previously were assured to only use information readily
available from the tuple header. This is considered acceptable, because
the extra I/O would involve cases that would previously cause some
commands to block waiting for concurrent transactions to finish.
Another important change is the fact that locking tuples that have
previously been updated causes the future versions to be marked as
locked, too; this is essential for correctness of foreign key checks.
This causes additional WAL-logging, also (there was previously a single
WAL record for a locked tuple; now there are as many as updated copies
of the tuple there exist.)
With all this in place, contention related to tuples being checked by
foreign key rules should be much reduced.
As a bonus, the old behavior that a subtransaction grabbing a stronger
tuple lock than the parent (sub)transaction held on a given tuple and
later aborting caused the weaker lock to be lost, has been fixed.
Many new spec files were added for isolation tester framework, to ensure
overall behavior is sane. There's probably room for several more tests.
There were several reviewers of this patch; in particular, Noah Misch
and Andres Freund spent considerable time in it. Original idea for the
patch came from Simon Riggs, after a problem report by Joel Jacobson.
Most code is from me, with contributions from Marti Raudsepp, Alexander
Shulgin, Noah Misch and Andres Freund.
This patch was discussed in several pgsql-hackers threads; the most
important start at the following message-ids:
AANLkTimo9XVcEzfiBR-ut3KVNDkjm2Vxh+t8kAmWjPuv@mail.gmail.com1290721684-sup-3951@alvh.no-ip.org1294953201-sup-2099@alvh.no-ip.org1320343602-sup-2290@alvh.no-ip.org1339690386-sup-8927@alvh.no-ip.org4FE5FF020200002500048A3D@gw.wicourts.gov4FEAB90A0200002500048B7D@gw.wicourts.gov
This got broken in the original fast-path locking patch, because
I failed to account for the fact that Hot Standby startup process
might take a strong relation lock on a relation in a database to
which it is not bound, and confused MyDatabaseId with the database
ID of the relation being locked.
Report and diagnosis by Andres Freund. Final form of patch by me.
When relations are dropped, at end of transaction we need to remove the
files and clean the buffer pool of buffers containing pages of those
relations. Previously we would scan the buffer pool once per relation
to clean up buffers. When there are many relations to drop, the
repeated scans make this process slow; so we now instead pass a list of
relations to drop and scan the pool once, checking each buffer against
the passed list. When the number of relations is larger than a
threshold (which as of this patch is being set to 20 relations) we sort
the array before starting, and bsearch the array; when it's smaller, we
simply scan the array linearly each time, because that's faster. The
exact optimal threshold value depends on many factors, but the
difference is not likely to be significant enough to justify making it
user-settable.
This has been measured to be a significant win (a 15x win when dropping
100,000 relations; an extreme case, but reportedly a real one).
Author: Tomas Vondra, some tweaks by me
Reviewed by: Robert Haas, Shigeru Hanada, Andres Freund, Álvaro Herrera
The code in PostPrepare_Locks supposed that it could reassign locks to
the prepared transaction's dummy PGPROC by deleting the PROCLOCK table
entries and immediately creating new ones. This was safe when that code
was written, but since we invented partitioning of the shared lock table,
it's not safe --- another process could steal away the PROCLOCK entry in
the short interval when it's on the freelist. Then, if we were otherwise
out of shared memory, PostPrepare_Locks would have to PANIC, since it's
too late to back out of the PREPARE at that point.
Fix by inventing a dynahash.c function to atomically update a hashtable
entry's key. (This might possibly have other uses in future.)
This is an ancient bug that in principle we ought to back-patch, but the
odds of someone hitting it in the field seem really tiny, because (a) the
risk window is small, and (b) nobody runs servers with maxed-out lock
tables for long, because they'll be getting non-PANIC out-of-memory errors
anyway. So fixing it in HEAD seems sufficient, at least until the new
code has gotten some testing.
If VirtualXactLock() has to wait for a transaction that holds its VXID lock
as a fast-path lock, it must first convert the fast-path lock to a regular
lock. It failed to take the required "partition" lock on the main
shared-memory lock table while doing so. This is the direct cause of the
assert failure in GetLockStatusData() recently observed in the buildfarm,
but more worryingly it could result in arbitrary corruption of the shared
lock table if some other process were concurrently engaged in modifying the
same partition of the lock table. Fortunately, VirtualXactLock() is only
used by CREATE INDEX CONCURRENTLY and DROP INDEX CONCURRENTLY, so the
opportunities for failure are fewer than they might have been.
In passing, improve some comments and be a bit more consistent about
order of operations.
In situations where there are over 8MB of empty pages at the end of
a table, the truncation work for trailing empty pages takes longer
than deadlock_timeout, and there is frequent access to the table by
processes other than autovacuum, there was a problem with the
autovacuum worker process being canceled by the deadlock checking
code. The truncation work done by autovacuum up that point was
lost, and the attempt tried again by a later autovacuum worker. The
attempts could continue indefinitely without making progress,
consuming resources and blocking other processes for up to
deadlock_timeout each time.
This patch has the autovacuum worker checking whether it is
blocking any other thread at 20ms intervals. If such a condition
develops, the autovacuum worker will persist the work it has done
so far, release its lock on the table, and sleep in 50ms intervals
for up to 5 seconds, hoping to be able to re-acquire the lock and
try again. If it is unable to get the lock in that time, it moves
on and a worker will try to continue later from the point this one
left off.
While this patch doesn't change the rules about when and what to
truncate, it does cause the truncation to occur sooner, with less
blocking, and with the consumption of fewer resources when there is
contention for the table's lock.
The only user-visible change other than improved performance is
that the table size during truncation may change incrementally
instead of just once.
This problem exists in all supported versions but is infrequently
reported, although some reports of performance problems when
autovacuum runs might be caused by this. Initial commit is just the
master branch, but this should probably be backpatched once the
build farm and general developer usage confirm that there are no
surprising effects.
Jan Wieck
Background workers are postmaster subprocesses that run arbitrary
user-specified code. They can request shared memory access as well as
backend database connections; or they can just use plain libpq frontend
database connections.
Modules listed in shared_preload_libraries can register background
workers in their _PG_init() function; this is early enough that it's not
necessary to provide an extra GUC option, because the necessary extra
resources can be allocated early on. Modules can install more than one
bgworker, if necessary.
Care is taken that these extra processes do not interfere with other
postmaster tasks: only one such process is started on each ServerLoop
iteration. This means a large number of them could be waiting to be
started up and postmaster is still able to quickly service external
connection requests. Also, shutdown sequence should not be impacted by
a worker process that's reasonably well behaved (i.e. promptly responds
to termination signals.)
The current implementation lets worker processes specify their start
time, i.e. at what point in the server startup process they are to be
started: right after postmaster start (in which case they mustn't ask
for shared memory access), when consistent state has been reached
(useful during recovery in a HOT standby server), or when recovery has
terminated (i.e. when normal backends are allowed).
In case of a bgworker crash, actions to take depend on registration
data: if shared memory was requested, then all other connections are
taken down (as well as other bgworkers), just like it were a regular
backend crashing. The bgworker itself is restarted, too, within a
configurable timeframe (which can be configured to be never).
More features to add to this framework can be imagined without much
effort, and have been discussed, but this seems good enough as a useful
unit already.
An elementary sample module is supplied.
Author: Álvaro Herrera
This patch is loosely based on prior patches submitted by KaiGai Kohei,
and unsubmitted code by Simon Riggs.
Reviewed by: KaiGai Kohei, Markus Wanner, Andres Freund,
Heikki Linnakangas, Simon Riggs, Amit Kapila
Rename PGXACT->inCommit flag into delayChkpt flag,
and generalise comments to allow use in other situations,
such as the forthcoming potential use in checksum patch.
Replace wait loop to look for VXIDs with delayChkpt set.
No user visible changes, not behaviour changes at present.
Simon Riggs, reviewed and rebased by Jeff Davis
This reverts commit c11130690d in favor of
actually fixing the problem: namely, that we should never have been
modifying the checkpoint record's nextXid at this point to begin with.
The nextXid should match the state as of the checkpoint's logical WAL
position (ie the redo point), not the state as of its physical position.
It's especially bogus to advance it in some wal_levels and not others.
In any case there is no need for the checkpoint record to carry the
same nextXid shown in the XLOG_RUNNING_XACTS record just emitted by
LogStandbySnapshot, as any replay operation will already have adopted
that value as current.
This fixes bug #7710 from Tarvi Pillessaar, and probably also explains bug
#6291 from Daniel Farina, in that if a checkpoint were in progress at the
instant of XID wraparound, the epoch bump would be lost as reported.
(And, of course, these days there's at least a 50-50 chance of a checkpoint
being in progress at any given instant.)
Diagnosed by me and independently by Andres Freund. Back-patch to all
branches supporting hot standby.
Previously we stored all xids mixed together.
Now we store top-level xids first, followed
by all subxids. Also skip logging any subxids
if the snapshot is suboverflowed, since there
are potentially large numbers of them and they
are not useful in that case anyway. Has value
in the envisaged design for decoding of WAL.
No planned effect on Hot Standby.
Andres Freund, reviewed by me
Files opened with BasicOpenFile or PathNameOpenFile are not automatically
cleaned up on error. That puts unnecessary burden on callers that only want
to keep the file open for a short time. There is AllocateFile, but that
returns a buffered FILE * stream, which in many cases is not the nicest API
to work with. So add function called OpenTransientFile, which returns a
unbuffered fd that's cleaned up like the FILE* returned by AllocateFile().
This plugs a few rare fd leaks in error cases:
1. copy_file() - fixed by by using OpenTransientFile instead of BasicOpenFile
2. XLogFileInit() - fixed by adding close() calls to the error cases. Can't
use OpenTransientFile here because the fd is supposed to persist over
transaction boundaries.
3. lo_import/lo_export - fixed by using OpenTransientFile instead of
PathNameOpenFile.
In addition to plugging those leaks, this replaces many BasicOpenFile() calls
with OpenTransientFile() that were not leaking, because the code meticulously
closed the file on error. That wasn't strictly necessary, but IMHO it's good
for robustness.
The same leaks exist in older versions, but given the rarity of the issues,
I'm not backpatching this. Not yet, anyway - it might be good to backpatch
later, after this mechanism has had some more testing in master branch.
errcontext() is typically used in an error context callback function, not
within an ereport() invocation like e.g errmsg and errdetail are. That means
that the message domain that the TEXTDOMAIN magic in ereport() determines
is not the right one for the errcontext() calls. The message domain needs to
be determined by the C file containing the errcontext() call, not the file
containing the ereport() call.
Fix by turning errcontext() into a macro that passes the TEXTDOMAIN to use
for the errcontext message. "errcontext" was used in a few places as a
variable or struct field name, I had to rename those out of the way, now
that errcontext is a macro.
We've had this problem all along, but this isn't doesn't seem worth
backporting. It's a fairly minor issue, and turning errcontext from a
function to a macro requires at least a recompile of any external code that
calls errcontext().
If an SMgrRelation is not "owned" by a relcache entry, don't allow it to
live past transaction end. This design allows the same SMgrRelation to be
used for blind writes of multiple blocks during a transaction, but ensures
that we don't hold onto such an SMgrRelation indefinitely. Because an
SMgrRelation typically corresponds to open file descriptors at the fd.c
level, leaving it open when there's no corresponding relcache entry can
mean that we prevent the kernel from reclaiming deleted disk space.
(While CacheInvalidateSmgr messages usually fix that, there are cases
where they're not issued, such as DROP DATABASE. We might want to add
some more sinval messaging for that, but I'd be inclined to keep this
type of logic anyway, since allowing VFDs to accumulate indefinitely
for blind-written relations doesn't seem like a good idea.)
This code replaces a previous attempt towards the same goal that proved
to be unreliable. Back-patch to 9.1 where the previous patch was added.
This reverts commit fba105b109.
That approach had problems with the smgr-level state not tracking what
we really want to happen, and with the VFD-level state not tracking the
smgr-level state very well either. In consequence, it was still possible
to hold kernel file descriptors open for long-gone tables (as in recent
report from Tore Halset), and yet there were also cases of FDs being closed
undesirably soon. A replacement implementation will follow.
In the previous coding, new backend processes would attempt to create their
self-pipe during the OwnLatch call in InitProcess. However, pipe creation
could fail if the kernel is short of resources; and the system does not
recover gracefully from a FATAL error right there, since we have armed the
dead-man switch for this process and not yet set up the on_shmem_exit
callback that would disarm it. The postmaster then forces an unnecessary
database-wide crash and restart, as reported by Sean Chittenden.
There are various ways we could rearrange the code to fix this, but the
simplest and sanest seems to be to split out creation of the self-pipe into
a new function InitializeLatchSupport, which must be called from a place
where failure is allowed. For most processes that gets called in
InitProcess or InitAuxiliaryProcess, but processes that don't call either
but still use latches need their own calls.
Back-patch to 9.1, which has only a part of the latch logic that 9.2 and
HEAD have, but nonetheless includes this bug.
Do read/write permissions checks at most once per large object descriptor,
not once per lo_read or lo_write call as before. The repeated tests were
quite useless in the read case since the snapshot-based tests were
guaranteed to produce the same answer every time. In the write case,
the extra tests could in principle detect revocation of write privileges
after a series of writes has started --- but there's a race condition there
anyway, since we'd check privileges before performing and certainly before
committing the write. So there's no real advantage to checking every
single time, and we might as well redefine it as "only check the first
time".
On the same reasoning, remove the LargeObjectExists checks in inv_write
and inv_truncate. We already checked existence when the descriptor was
opened, and checking again doesn't provide any real increment of safety
that would justify the cost.
