Commit 45f6240a8f added an assumption in ExecHashIncreaseNumBatches
and ExecHashIncreaseNumBuckets that they could find all tuples in the main
hash table by iterating over the "dense storage" introduced by that patch.
However, ExecHashRemoveNextSkewBucket continued its old practice of simply
re-linking deleted skew tuples into the main table's hashchains. Hence,
such tuples got lost during any subsequent increase in nbatch or nbuckets,
and would never get joined, as reported in bug #13908 from Seth P.
I (tgl) think that the aforesaid commit has got multiple design issues
and should be reworked rather completely; but there is no time for that
right now, so band-aid the problem by making ExecHashRemoveNextSkewBucket
physically copy deleted skew tuples into the "dense storage" arena.
The added test case is able to exhibit the problem by means of fooling the
planner with a WHERE condition that it will underestimate the selectivity
of, causing the initial nbatch estimate to be too small.
Tomas Vondra and Tom Lane. Thanks to David Johnston for initial
investigation into the bug report.
Commit 30d7ae3c76 introduced an HJDEBUG
stanza that probably didn't compile at the time, and definitely doesn't
compile now, because it refers to a nonexistent variable. It doesn't seem
terribly useful anyway, so just get rid of it.
While I'm fooling with it, use %z modifier instead of the obsolete hack of
casting size_t to unsigned long, and include the HashJoinTable's address in
each printout so that it's possible to distinguish the activities of
multiple hashjoins occurring in one query.
Noted while trying to use HJDEBUG to investigate bug #13908. Back-patch
to 9.5, because code that doesn't compile is certainly not very helpful.
This patch doesn't put the new infrastructure to use anywhere, and
indeed it's not clear how it could ever be used for something like
postgres_fdw which has to send an SQL query and wait for a reply,
but there might be FDWs or custom scan providers that are CPU-bound,
so let's give them a way to join club parallel.
KaiGai Kohei, reviewed by me.
Previously, the foreign join pushdown infrastructure left the question
of security entirely up to individual FDWs, but it would be easy for
a foreign data wrapper to inadvertently open up subtle security holes
that way. So, make it the core code's job to determine which user
mapping OID is relevant, and don't attempt join pushdown unless it's
the same for all relevant relations.
Per a suggestion from Tom Lane. Shigeru Hanada and Ashutosh Bapat,
reviewed by Etsuro Fujita and KaiGai Kohei, with some further
changes by me.
The core innovation of this patch is the introduction of the concept
of a partial path; that is, a path which if executed in parallel will
generate a subset of the output rows in each process. Gathering a
partial path produces an ordinary (complete) path. This allows us to
generate paths for parallel joins by joining a partial path for one
side (which at the baserel level is currently always a Partial Seq
Scan) to an ordinary path on the other side. This is subject to
various restrictions at present, especially that this strategy seems
unlikely to be sensible for merge joins, so only nested loops and
hash joins paths are generated.
This also allows an Append node to be pushed below a Gather node in
the case of a partitioned table.
Testing revealed that early versions of this patch made poor decisions
in some cases, which turned out to be caused by the fact that the
original cost model for Parallel Seq Scan wasn't very good. So this
patch tries to make some modest improvements in that area.
There is much more to be done in the area of generating good parallel
plans in all cases, but this seems like a useful step forward.
Patch by me, reviewed by Dilip Kumar and Amit Kapila.
Aggregate nodes now have two new modes: a "partial" mode where they
output the unfinalized transition state, and a "finalize" mode where
they accept unfinalized transition states rather than individual
values as input.
These new modes are not used anywhere yet, but they will be necessary
for parallel aggregation. The infrastructure also figures to be
useful for cases where we want to aggregate local data and remote
data via the FDW interface, and want to bring back partial aggregates
from the remote side that can then be combined with locally generated
partial aggregates to produce the final value. It may also be useful
even when neither FDWs nor parallelism are in play, as explained in
the comments in nodeAgg.c.
David Rowley and Simon Riggs, reviewed by KaiGai Kohei, Heikki
Linnakangas, Haribabu Kommi, and me.
This patch reduces pg_am to just two columns, a name and a handler
function. All the data formerly obtained from pg_am is now provided
in a C struct returned by the handler function. This is similar to
the designs we've adopted for FDWs and tablesample methods. There
are multiple advantages. For one, the index AM's support functions
are now simple C functions, making them faster to call and much less
error-prone, since the C compiler can now check function signatures.
For another, this will make it far more practical to define index access
methods in installable extensions.
A disadvantage is that SQL-level code can no longer see attributes
of index AMs; in particular, some of the crosschecks in the opr_sanity
regression test are no longer possible from SQL. We've addressed that
by adding a facility for the index AM to perform such checks instead.
(Much more could be done in that line, but for now we're content if the
amvalidate functions more or less replace what opr_sanity used to do.)
We might also want to expose some sort of reporting functionality, but
this patch doesn't do that.
Alexander Korotkov, reviewed by Petr Jelínek, and rather heavily
editorialized on by me.
The original coding read tuples from workers in round-robin fashion,
but performance testing shows that it works much better to read enough
to empty one queue before moving on to the next. I believe the
reason for this is that, with the old approach, we could easily wake
up a worker repeatedly to write only one new tuple into the shm_mq
each time. With this approach, by the time the process gets scheduled,
it has a decent chance of being able to fill the entire buffer in
one go.
Patch by me. Dilip Kumar helped with performance testing.
Omitted boundaries represent the upper or lower limit of the corresponding
array subscript. This allows simpler specification of many common
use-cases.
(Revised version of commit 9246af6799)
YUriy Zhuravlev
ExecOnConflictUpdate() passed t_ctid of the to-be-updated tuple to
ExecUpdate(). That's problematic primarily because of two reason: First
and foremost t_ctid could point to a different tuple. Secondly, and
that's what triggered the complaint by Stanislav, t_ctid is changed by
heap_update() to point to the new tuple version. The behavior of AFTER
UPDATE triggers was therefore broken, with NEW.* and OLD.* tuples
spuriously identical within AFTER UPDATE triggers.
To fix both issues, pass a pointer to t_self of a on-stack HeapTuple
instead.
Fixing this bug lead to one change in regression tests, which previously
failed due to the first issue mentioned above. There's a reasonable
expectation that test fails, as it updates one row repeatedly within one
INSERT ... ON CONFLICT statement. That is only possible if the second
update is triggered via ON CONFLICT ... SET, ON CONFLICT ... WHERE, or
by a WITH CHECK expression, as those are executed after
ExecOnConflictUpdate() does a visibility check. That could easily be
prohibited, but given it's allowed for plain UPDATEs and a rare corner
case, it doesn't seem worthwhile.
Reported-By: Stanislav Grozev
Author: Andres Freund and Peter Geoghegan
Discussion: CAA78GVqy1+LisN-8DygekD_Ldfy=BJLarSpjGhytOsgkpMavfQ@mail.gmail.com
Backpatch: 9.5, where ON CONFLICT was introduced
The original parallel sequential scan commit included only very limited
changes to the EXPLAIN output. Aggregated totals from all workers were
displayed, but there was no way to see what each individual worker did
or to distinguish the effort made by the workers from the effort made by
the leader.
Per a gripe by Thom Brown (and maybe others). Patch by me, reviewed
by Amit Kapila.
Commit e7cb7ee145 provided basic
infrastructure for allowing a foreign data wrapper or custom scan
provider to replace a join of one or more tables with a scan.
However, this infrastructure failed to take into account the need
for possible EvalPlanQual rechecks, and ExecScanFetch would fail
an assertion (or just overwrite memory) if such a check was attempted
for a plan containing a pushed-down join. To fix, adjust the EPQ
machinery to skip some processing steps when scanrelid == 0, making
those the responsibility of scan's recheck method, which also has
the responsibility in this case of correctly populating the relevant
slot.
To allow foreign scans to gain control in the right place to make
use of this new facility, add a new, optional RecheckForeignScan
method. Also, allow a foreign scan to have a child plan, which can
be used to correctly populate the slot (or perhaps for something
else, but this is the only use currently envisioned).
KaiGai Kohei, reviewed by Robert Haas, Etsuro Fujita, and Kyotaro
Horiguchi.
The previous coding attempts to destroy the DSM in this case, but
child nodes might have stored data there and still be holding onto
pointers in this case. So don't do that.
Also, free the reader array instead of leaking it.
Extracted from two different patch versions both by Amit Kapila.
When I wrote this code originally, the intention was to recompute the
remapinfo only when the tupledesc changes. This presumably only
happens once per query, but I copied the design pattern from other
DestReceivers. However, due to a silly oversight on my part,
tqueue->tupledesc never got set, leading to recomputation for every
tuple.
This should improve the performance of parallel scans that return a
significant number of tuples.
Report by Amit Kapila; patch by me, reviewed by him.
In addition, this path fills in a number of missing bits and pieces in
the parallel infrastructure. Paths and plans now have a parallel_aware
flag indicating whether whatever parallel-aware logic they have should
be engaged. It is believed that we will need this flag for a number of
path/plan types, not just sequential scans, which is why the flag is
generic rather than part of the SeqScan structures specifically.
Also, execParallel.c now gives parallel nodes a chance to initialize
their PlanState nodes from the DSM during parallel worker startup.
Amit Kapila, with a fair amount of adjustment by me. Review of previous
patch versions by Haribabu Kommi and others.
The patches committed as 6e71dd7ce9
and 3a1f8611f2 were developed in
parallel but dependent on each other in a way that I failed to
notice.
This patch to fix the problem was prepared by Amit Kapila.
Commit 4a4e6893aa, which introduced this
mechanism, failed to account for the fact that the RECORD pseudo-type
uses transient typmods that are only meaningful within a single
backend. Transferring such tuples without modification between two
cooperating backends does not work. This commit installs a system
for passing the tuple descriptors over the same shm_mq being used to
send the tuples themselves. The two sides might not assign the same
transient typmod to any given tuple descriptor, so we must also
substitute the appropriate receiver-side typmod for the one used by
the sender. That adds some CPU overhead, but still seems better than
being unable to pass records between cooperating parallel processes.
Along the way, move the logic for handling multiple tuple queues from
tqueue.c to nodeGather.c; tqueue.c now provides a TupleQueueReader,
which reads from a single queue, rather than a TupleQueueFunnel, which
potentially reads from multiple queues. This change was suggested
previously as a way to make sure that nodeGather.c rather than tqueue.c
had policy control over the order in which to read from queues, but
it wasn't clear to me until now how good an idea it was. typmod
mapping needs to be performed separately for each queue, and it is
much simpler if the tqueue.c code handles that and leaves multiplexing
multiple queues to higher layers of the stack.
Commit d1b7c1ffe7 introduced a mechanism
for serializing a ParamListInfo structure to be passed to a parallel
worker. However, this mechanism failed to handle external expanded
values, as pointed out by Noah Misch. Repair.
Moreover, plpgsql_param_fetch requires adjustment because the
serialization mechanism needs it to skip evaluating unused parameters
just as we would do when it is called from copyParamList, but params
== estate->paramLI in that case. To fix, make the bms_is_member test
in that function unconditional.
Finally, have setup_param_list set a new ParamListInfo field,
paramMask, to the parameters actually used in the expression, so that
we don't try to fetch those that are not needed when serializing a
parameter list. This isn't necessary for correctness, but it makes
the performance of the parallel executor code comparable to what we
do for cases involving cursors.
Design suggestions and extensive review by Noah Misch. Patch by me.
Commit b0b0d84b3d purported to make it
possible to relaunch workers using the same parallel context, but it had
an unpleasant race condition: we might reinitialize after the workers
have sent their last control message but before they have dettached the
DSM, leaving to crashes. Repair by introducing a new ParallelContext
operation, ReinitializeParallelDSM.
Adjust execParallel.c to use this new support, so that we can rescan a
Gather node by relaunching workers but without needing to recreate the
DSM.
Amit Kapila, with some adjustments by me. Extracted from latest parallel
sequential scan patch.
The original Gather code failed to mark a Gather node as not able to
do projection, but it couldn't, even though it did call initialize its
projection info via ExecAssignProjectionInfo. There doesn't seem to
be any good reason for this node not to have projection capability,
so clean things up so that it does. Without this, plans using Gather
nodes might need to carry extra Result nodes to do projection.
Commit 816e336f12 added the wrong error
check to async.c; sending restrictions is restricted to the leader,
not altogether unsafe.
Commit 3bd909b220 added ExecShutdownNode
to traverse the planstate tree and call shutdown functions, but made
a Gather node, the only node that actually has such a function, abort
the tree traversal, which is wrong.
In the previous coding, before returning from ExecutorRun, we'd shut
down all parallel workers. This was dead wrong if ExecutorRun was
called with a non-zero tuple count; it had the effect of truncating
the query output. To fix, give ExecutePlan control over whether to
enter parallel mode, and have it refuse to do so if the tuple count
is non-zero. Rewrite the Gather logic so that it can cope with being
called outside parallel mode.
Commit 7aea8e4f2d is largely to blame
for this problem, though this patch modifies some subsequently-committed
code which relied on the guarantees it purported to make.
This fixes a long-standing bug which was discovered while investigating
the interaction between the new join pushdown code and the EvalPlanQual
machinery: if a ForeignScan appears on the inner side of a paramaterized
nestloop, an EPQ recheck would re-return the original tuple even if
it no longer satisfied the pushed-down quals due to changed parameter
values.
This fix adds a new member to ForeignScan and ForeignScanState and a
new argument to make_foreignscan, and requires changes to FDWs which
push down quals to populate that new argument with a list of quals they
have chosen to push down. Therefore, I'm only back-patching to 9.5,
even though the bug is not new in 9.5.
Etsuro Fujita, reviewed by me and by Kyotaro Horiguchi.
On reflection, the submitted patch didn't really work to prevent the
request size from exceeding MaxAllocSize, because of the fact that we'd
happily round nbuckets up to the next power of 2 after we'd limited it to
max_pointers. The simplest way to enforce the limit correctly is to
round max_pointers down to a power of 2 when it isn't one already.
(Note that the constraint to INT_MAX / 2, if it were doing anything useful
at all, is properly applied after that.)
Limit the size of the hashtable pointer array to not more than
MaxAllocSize, per reports from Kouhei Kaigai and others of "invalid memory
alloc request size" failures. There was discussion of allowing the array
to get larger than that by using the "huge" palloc API, but so far no proof
that that is actually a good idea, and at this point in the 9.5 cycle major
changes from old behavior don't seem like the way to go.
Fix a rather serious secondary bug in the new code, which was that it
didn't ensure nbuckets remained a power of 2 when recomputing it for the
multiple-batch case.
Clean up sloppy division of labor between ExecHashIncreaseNumBuckets and
its sole call site.
A Gather executor node runs any number of copies of a plan in an equal
number of workers and merges all of the results into a single tuple
stream. It can also run the plan itself, if the workers are
unavailable or haven't started up yet. It is intended to work with
the Partial Seq Scan node which will be added in future commits.
It could also be used to implement parallel query of a different sort
by itself, without help from Partial Seq Scan, if the single_copy mode
is used. In that mode, a worker executes the plan, and the parallel
leader does not, merely collecting the worker's results. So, a Gather
node could be inserted into a plan to split the execution of that plan
across two processes. Nested Gather nodes aren't currently supported,
but we might want to add support for that in the future.
There's nothing in the planner to actually generate Gather nodes yet,
so it's not quite time to break out the champagne. But we're getting
close.
Amit Kapila. Some designs suggestions were provided by me, and I also
reviewed the patch. Single-copy mode, documentation, and other minor
changes also by me.
This code provides infrastructure for a parallel leader to start up
parallel workers to execute subtrees of the plan tree being executed
in the master. User-supplied parameters from ParamListInfo are passed
down, but PARAM_EXEC parameters are not. Various other constructs,
such as initplans, subplans, and CTEs, are also not currently shared.
Nevertheless, there's enough here to support a basic implementation of
parallel query, and we can lift some of the current restrictions as
needed.
Amit Kapila and Robert Haas
When taking the UPDATE path in an INSERT .. ON CONFLICT .. UPDATE tables
with oids were not supported. The tuple generated by the update target
list was projected without space for an oid - a simple oversight.
Reported-By: Peter Geoghegan
Author: Andres Freund
Backpatch: 9.5, where ON CONFLICT was introduced
The shm_mq mechanism was built to send error (and notice) messages and
tuples between backends. However, shm_mq itself only deals in raw
bytes. Since commit 2bd9e412f9, we have
had infrastructure for one message to redirect protocol messages to a
queue and for another backend to parse them and do useful things with
them. This commit introduces a somewhat analogous facility for tuples
by adding a new type of DestReceiver, DestTupleQueue, which writes
each tuple generated by a query into a shm_mq, and a new
TupleQueueFunnel facility which reads raw tuples out of the queue and
reconstructs the HeapTuple format expected by the executor.
The TupleQueueFunnel abstraction supports reading from multiple tuple
streams at the same time, but only in round-robin fashion. Someone
could imaginably want other policies, but this should be good enough
to meet our short-term needs related to parallel query, and we can
always extend it later.
This also makes one minor addition to the shm_mq API that didn'
seem worth breaking out as a separate patch.
Extracted from Amit Kapila's parallel sequential scan patch. This
code was originally written by me, and then it was revised by Amit,
and then it was revised some more by me.
Commit 924bcf4f16 introduced a framework
for parallel computation in PostgreSQL that makes most but not all
built-in functions safe to execute in parallel mode. In order to have
parallel query, we'll need to be able to determine whether that query
contains functions (either built-in or user-defined) that cannot be
safely executed in parallel mode. This requires those functions to be
labeled, so this patch introduces an infrastructure for that. Some
functions currently labeled as safe may need to be revised depending on
how pending issues related to heavyweight locking under paralllelism
are resolved.
Parallel plans can't be used except for the case where the query will
run to completion. If portal execution were suspended, the parallel
mode restrictions would need to remain in effect during that time, but
that might make other queries fail. Therefore, this patch introduces
a framework that enables consideration of parallel plans only when it
is known that the plan will be run to completion. This probably needs
some refinement; for example, at bind time, we do not know whether a
query run via the extended protocol will be execution to completion or
run with a limited fetch count. Having the client indicate its
intentions at bind time would constitute a wire protocol break. Some
contexts in which parallel mode would be safe are not adjusted by this
patch; the default is not to try parallel plans except from call sites
that have been updated to say that such plans are OK.
This commit doesn't introduce any parallel paths or plans; it just
provides a way to determine whether they could potentially be used.
I'm committing it on the theory that the remaining parallel sequential
scan patches will also get committed to this release, hopefully in the
not-too-distant future.
Robert Haas and Amit Kapila. Reviewed (in earlier versions) by Noah
Misch.
This refactors rewrite/rowsecurity.c to simplify the handling of the
default deny case (reducing the number of places where we check for and
add the default deny policy from three to one) by splitting up the
retrival of the policies from the application of them.
This also allowed us to do away with the policy_id field. A policy_name
field was added for WithCheckOption policies and is used in error
reporting, when available.
Patch by Dean Rasheed, with various mostly cosmetic changes by me.
Back-patch to 9.5 where RLS was introduced to avoid unnecessary
differences, since we're still in alpha, per discussion with Robert.
Per discussion, nowadays it is possible to have tablespaces that have
wildly different I/O characteristics from others. Setting different
effective_io_concurrency parameters for those has been measured to
improve performance.
Author: Julien Rouhaud
Reviewed by: Andres Freund
For no obvious reason, spi_printtup() was coded to enlarge the tuple
pointer table by just 256 slots at a time, rather than doubling the size at
each reallocation, as is our usual habit. For very large SPI results, this
makes for O(N^2) time spent in repalloc(), which of course soon comes to
dominate the runtime. Use the standard doubling approach instead.
This is a longstanding performance bug, so back-patch to all active
branches.
Neil Conway
Commit 8cce08f168 used a left-shift
on a literal of 1 that could (in large allocations) be shifted by
31 or more bits. This was assigned to a local variable that was
already declared to be a long to protect against overruns of int,
but the literal in this shift needs to be declared long to allow it
to work correctly in some compilers.
Backpatch to 9.5, where the bug was introduced.
Report and patch by KaiGai Kohei, slighly modified based on
discussion.
It was a bool, even though it should be CEOUC_WAIT_MODE. That's unlikely
to have a negative effect with the current definition of bool (char),
but it's definitely wrong.
Discussion: 20150812084351.GD8470@awork2.anarazel.de
Backpatch: 9.5, where ON CONFLICT was merged
If there are two different aggregates in the query with same inputs, and
the aggregates have the same initial condition and transition function,
only calculate the state value once, and only call the final functions
separately. For example, AVG(x) and SUM(x) aggregates have the same
transition function, which accumulates the sum and number of input tuples.
For a query like "SELECT AVG(x), SUM(x) FROM x", we can therefore
accumulate the state function only once, which gives a nice speedup.
David Rowley, reviewed and edited by me.
It's against project policy to use elog() for user-facing errors, or to
omit an errcode() selection for errors that aren't supposed to be "can't
happen" cases. Fix all the violations of this policy that result in
ERRCODE_INTERNAL_ERROR log entries during the standard regression tests,
as errors that can reliably be triggered from SQL surely should be
considered user-facing.
I also looked through all the files touched by this commit and fixed
other nearby problems of the same ilk. I do not claim to have fixed
all violations of the policy, just the ones in these files.
In a few places I also changed existing ERRCODE choices that didn't
seem particularly appropriate; mainly replacing ERRCODE_SYNTAX_ERROR
by something more specific.
Back-patch to 9.5, but no further; changing ERRCODE assignments in
stable branches doesn't seem like a good idea.
Although I think on all modern machines floating division by zero
results in Infinity not SIGFPE, we still don't want infinities
running around in the planner's costing estimates; too much risk
of that leading to insane behavior.
grouping_planner() failed to consider the possibility that final_rel
might be known dummy and hence have zero rowcount. (I wonder if it
would be better to set a rows estimate of 1 for dummy relations?
But at least in the back branches, changing this convention seems
like a bad idea, so I'll leave that for another day.)
Make certain that get_variable_numdistinct() produces a nonzero result.
The case that can be shown to be broken is with stadistinct < 0.0 and
small ntuples; we did not prevent the result from rounding to zero.
For good luck I applied clamp_row_est() to all the nonconstant return
values.
In ExecChooseHashTableSize(), Assert that we compute positive nbuckets
and nbatch. I know of no reason to think this isn't the case, but it
seems like a good safety check.
Per reports from Piotr Stefaniak. Back-patch to all active branches.
The pg_stats view is supposed to be restricted to only show rows
about tables the user can read. However, it sometimes can leak
information which could not otherwise be seen when row level security
is enabled. Fix that by not showing pg_stats rows to users that would
be subject to RLS on the table the row is related to. This is done
by creating/using the newly introduced SQL visible function,
row_security_active().
Along the way, clean up three call sites of check_enable_rls(). The second
argument of that function should only be specified as other than
InvalidOid when we are checking as a different user than the current one,
as in when querying through a view. These sites were passing GetUserId()
instead of InvalidOid, which can cause the function to return incorrect
results if the current user has the BYPASSRLS privilege and row_security
has been set to OFF.
Additionally fix a bug causing RI Trigger error messages to unintentionally
leak information when RLS is enabled, and other minor cleanup and
improvements. Also add WITH (security_barrier) to the definition of pg_stats.
Bumped CATVERSION due to new SQL functions and pg_stats view definition.
Back-patch to 9.5 where RLS was introduced. Reported by Yaroslav.
Patch by Joe Conway and Dean Rasheed with review and input by
Michael Paquier and Stephen Frost.
There is no full discussion of speculative insertions in the executor
README. There is a high-level explanation in execIndexing.c, but it doesn't
seem necessary to refer it from here.
Peter Geoghegan
The original implementation of TABLESAMPLE modeled the tablesample method
API on index access methods, which wasn't a good choice because, without
specialized DDL commands, there's no way to build an extension that can
implement a TSM. (Raw inserts into system catalogs are not an acceptable
thing to do, because we can't undo them during DROP EXTENSION, nor will
pg_upgrade behave sanely.) Instead adopt an API more like procedural
language handlers or foreign data wrappers, wherein the only SQL-level
support object needed is a single handler function identified by having
a special return type. This lets us get rid of the supporting catalog
altogether, so that no custom DDL support is needed for the feature.
Adjust the API so that it can support non-constant tablesample arguments
(the original coding assumed we could evaluate the argument expressions at
ExecInitSampleScan time, which is undesirable even if it weren't outright
unsafe), and discourage sampling methods from looking at invisible tuples.
Make sure that the BERNOULLI and SYSTEM methods are genuinely repeatable
within and across queries, as required by the SQL standard, and deal more
honestly with methods that can't support that requirement.
Make a full code-review pass over the tablesample additions, and fix
assorted bugs, omissions, infelicities, and cosmetic issues (such as
failure to put the added code stanzas in a consistent ordering).
Improve EXPLAIN's output of tablesample plans, too.
Back-patch to 9.5 so that we don't have to support the original API
in production.
Fix some places where pgindent did silly stuff, often because project
style wasn't followed to begin with. (I've not touched the atomics
headers, though.)
Fix confusion in documentation, substantial memory leakage if float8 or
float4 are pass-by-reference, and assorted comments that were obsoleted
by commit 98edd617f3.
Paul Ramsey reported that commit 35fcb1b3d0
induced a core dump on commuted ORDER BY expressions, because it was
assuming that the indexorderby expression could be found verbatim in the
relevant equivalence class, but it wasn't there. We really don't need
anything that complicated anyway; for the data types likely to be used for
index ORDER BY operators in the foreseeable future, the exprType() of the
ORDER BY expression will serve fine. (The case where we'd have to work
harder is where the ORDER BY expression's result is only binary-compatible
with the declared input type of the ordering operator; long before worrying
about that, one would need to get rid of GiST's hard-wired assumption that
said datatype is float8.)
Aside from fixing that crash and adding a regression test for the case,
I did some desultory code review:
nodeIndexscan.c was likewise overthinking how hard it ought to work to
identify the datatype of the ORDER BY expressions.
Add comments explaining how come nodeIndexscan.c can get away with
simplifying assumptions about NULLS LAST ordering and no backward scan.
Revert no-longer-needed changes of find_ec_member_for_tle(); while the
new definition was no worse than the old, it wasn't better either, and
it might cause back-patching pain.
Revert entirely bogus additions to genam.h.
Use "a" and "an" correctly, mostly in comments. Two error messages were
also fixed (they were just elogs, so no translation work required). Two
function comments in pg_proc.h were also fixed. Etsuro Fujita reported one
of these, but I found a lot more with grep.
Also fix a few other typos spotted while grepping for the a/an typos.
For example, "consists out of ..." -> "consists of ...". Plus a "though"/
"through" mixup reported by Euler Taveira.
Many of these typos were in old code, which would be nice to backpatch to
make future backpatching easier. But much of the code was new, and I didn't
feel like crafting separate patches for each branch. So no backpatching.
The previous coding was a leftover from attempting to hang all the on
conflict logic onto modify table's child nodes. It appears to not have
actually caused problems except for explain.
Add test exercising the broken and some other code paths.
Author: Peter Geoghegan and Andres Freund
This oversight results in a crash at executor startup if the plan has
been copied. outfuncs.c was missed as well.
While we could probably have taught both those files to cope with the
originally chosen representation of an Oid array, it would have been
painful, not least because there'd be no easy way to verify the array
length. An Oid List is far easier to work with. And AFAICS, there is
no particular notational benefit to using an array rather than a list
in the existing parts of the patch either. So just change it to a list.
Error in commit 35fcb1b3d0, which is new,
so no need for back-patch.
This SQL standard functionality allows to aggregate data by different
GROUP BY clauses at once. Each grouping set returns rows with columns
grouped by in other sets set to NULL.
This could previously be achieved by doing each grouping as a separate
query, conjoined by UNION ALLs. Besides being considerably more concise,
grouping sets will in many cases be faster, requiring only one scan over
the underlying data.
The current implementation of grouping sets only supports using sorting
for input. Individual sets that share a sort order are computed in one
pass. If there are sets that don't share a sort order, additional sort &
aggregation steps are performed. These additional passes are sourced by
the previous sort step; thus avoiding repeated scans of the source data.
The code is structured in a way that adding support for purely using
hash aggregation or a mix of hashing and sorting is possible. Sorting
was chosen to be supported first, as it is the most generic method of
implementation.
Instead of, as in an earlier versions of the patch, representing the
chain of sort and aggregation steps as full blown planner and executor
nodes, all but the first sort are performed inside the aggregation node
itself. This avoids the need to do some unusual gymnastics to handle
having to return aggregated and non-aggregated tuples from underlying
nodes, as well as having to shut down underlying nodes early to limit
memory usage. The optimizer still builds Sort/Agg node to describe each
phase, but they're not part of the plan tree, but instead additional
data for the aggregation node. They're a convenient and preexisting way
to describe aggregation and sorting. The first (and possibly only) sort
step is still performed as a separate execution step. That retains
similarity with existing group by plans, makes rescans fairly simple,
avoids very deep plans (leading to slow explains) and easily allows to
avoid the sorting step if the underlying data is sorted by other means.
A somewhat ugly side of this patch is having to deal with a grammar
ambiguity between the new CUBE keyword and the cube extension/functions
named cube (and rollup). To avoid breaking existing deployments of the
cube extension it has not been renamed, neither has cube been made a
reserved keyword. Instead precedence hacking is used to make GROUP BY
cube(..) refer to the CUBE grouping sets feature, and not the function
cube(). To actually group by a function cube(), unlikely as that might
be, the function name has to be quoted.
Needs a catversion bump because stored rules may change.
Author: Andrew Gierth and Atri Sharma, with contributions from Andres Freund
Reviewed-By: Andres Freund, Noah Misch, Tom Lane, Svenne Krap, Tomas
Vondra, Erik Rijkers, Marti Raudsepp, Pavel Stehule
Discussion: CAOeZVidmVRe2jU6aMk_5qkxnB7dfmPROzM7Ur8JPW5j8Y5X-Lw@mail.gmail.com
Add a TABLESAMPLE clause to SELECT statements that allows
user to specify random BERNOULLI sampling or block level
SYSTEM sampling. Implementation allows for extensible
sampling functions to be written, using a standard API.
Basic version follows SQLStandard exactly. Usable
concrete use cases for the sampling API follow in later
commits.
Petr Jelinek
Reviewed by Michael Paquier and Simon Riggs
We can only support a lossy distance function when the distance function's
datatype is comparable with the original ordering operator's datatype.
The distance function always returns a float8, so we are limited to float8,
and float4 (by a hard-coded cast of the float8 to float4).
In light of this limitation, it seems like a good idea to have a separate
'recheck' flag for the ORDER BY expressions, so that if you have a non-lossy
distance function, it still works with lossy quals. There are cases like
that with the build-in or contrib opclasses, but it's plausible.
There was a hidden assumption that the ORDER BY values returned by GiST
match the original ordering operator's return type, but there are plenty
of examples where that's not true, e.g. in btree_gist and pg_trgm. As long
as the distance function is not lossy, we can tolerate that and just not
return the distance to the executor (or rather, always return NULL). The
executor doesn't need the distances if there are no lossy results.
There was another little bug: the recheck variable was not initialized
before calling the distance function. That revealed the bigger issue,
as the executor tried to reorder tuples that didn't need reordering, and
that failed because of the datatype mismatch.
The distance function can now set *recheck = false, like index quals. The
executor will then re-check the ORDER BY expressions, and use a queue to
reorder the results on the fly.
This makes it possible to do kNN-searches on polygons and circles, which
don't store the exact value in the index, but just a bounding box.
Alexander Korotkov and me
This patch introduces the ability for complex datatypes to have an
in-memory representation that is different from their on-disk format.
On-disk formats are typically optimized for minimal size, and in any case
they can't contain pointers, so they are often not well-suited for
computation. Now a datatype can invent an "expanded" in-memory format
that is better suited for its operations, and then pass that around among
the C functions that operate on the datatype. There are also provisions
(rudimentary as yet) to allow an expanded object to be modified in-place
under suitable conditions, so that operations like assignment to an element
of an array need not involve copying the entire array.
The initial application for this feature is arrays, but it is not hard
to foresee using it for other container types like JSON, XML and hstore.
I have hopes that it will be useful to PostGIS as well.
In this initial implementation, a few heuristics have been hard-wired
into plpgsql to improve performance for arrays that are stored in
plpgsql variables. We would like to generalize those hacks so that
other datatypes can obtain similar improvements, but figuring out some
appropriate APIs is left as a task for future work. (The heuristics
themselves are probably not optimal yet, either, as they sometimes
force expansion of arrays that would be better left alone.)
Preliminary performance testing shows impressive speed gains for plpgsql
functions that do element-by-element access or update of large arrays.
There are other cases that get a little slower, as a result of added array
format conversions; but we can hope to improve anything that's annoyingly
bad. In any case most applications should see a net win.
Tom Lane, reviewed by Andres Freund
If a postgres_fdw foreign table is a non-locked source relation in an
UPDATE, DELETE, or SELECT FOR UPDATE/SHARE, and the query selects its
ctid column, the wrong value would be returned if an EvalPlanQual
recheck occurred. This happened because the foreign table's result row
was copied via the ROW_MARK_COPY code path, and EvalPlanQualFetchRowMarks
just unconditionally set the reconstructed tuple's t_self to "invalid".
To fix that, we can have EvalPlanQualFetchRowMarks copy the composite
datum's t_ctid field, and be sure to initialize that along with t_self
when postgres_fdw constructs a tuple to return.
If we just did that much then EvalPlanQualFetchRowMarks would start
returning "(0,0)" as ctid for all other ROW_MARK_COPY cases, which perhaps
does not matter much, but then again maybe it might. The cause of that is
that heap_form_tuple, which is the ultimate source of all composite datums,
simply leaves t_ctid as zeroes in newly constructed tuples. That seems
like a bad idea on general principles: a field that's really not been
initialized shouldn't appear to have a valid value. So let's eat the
trivial additional overhead of doing "ItemPointerSetInvalid(&(td->t_ctid))"
in heap_form_tuple.
This closes out our handling of Etsuro Fujita's report that tableoid and
ctid weren't correctly set in postgres_fdw EvalPlanQual cases. Along the
way we did a great deal of work to improve FDWs' ability to control row
locking behavior; which was not wasted effort by any means, but it didn't
end up being a fix for this problem because that feature would be too
expensive for postgres_fdw to use all the time.
Although the fix for the tableoid misbehavior was back-patched, I'm
hesitant to do so here; it seems far less likely that people would care
about remote ctid than tableoid, and even such a minor behavioral change
as this in heap_form_tuple is perhaps best not back-patched. So commit
to HEAD only, at least for the moment.
Etsuro Fujita, with some adjustments by me
Specifically the tlist and rti of the pseudo "excluded" relation weren't
properly treated by expression_tree_walker, which lead to errors when
excluded was referenced inside a rule because the varnos where not
properly adjusted. Similar omissions in OffsetVarNodes and
expression_tree_mutator had less impact, but should obviously be fixed
nonetheless.
A couple tests of for ON CONFLICT UPDATE into INSERT rule bearing
relations have been added.
In passing I updated a couple comments.
Previously, FDWs could only do "early row locking", that is lock a row as
soon as it's fetched, even though local restriction/join conditions might
discard the row later. This patch adds callbacks that allow FDWs to do
late locking in the same way that it's done for regular tables.
To make use of this feature, an FDW must support the "ctid" column as a
unique row identifier. Currently, since ctid has to be of type TID,
the feature is of limited use, though in principle it could be used by
postgres_fdw. We may eventually allow FDWs to specify another data type
for ctid, which would make it possible for more FDWs to use this feature.
This commit does not modify postgres_fdw to use late locking. We've
tested some prototype code for that, but it's not in committable shape,
and besides it's quite unclear whether it actually makes sense to do late
locking against a remote server. The extra round trips required are likely
to outweigh any benefit from improved concurrency.
Etsuro Fujita, reviewed by Ashutosh Bapat, and hacked up a lot by me
Commit e7cb7ee145 included some design
decisions that seem pretty questionable to me, and there was quite a lot
of stuff not to like about the documentation and comments. Clean up
as follows:
* Consider foreign joins only between foreign tables on the same server,
rather than between any two foreign tables with the same underlying FDW
handler function. In most if not all cases, the FDW would simply have had
to apply the same-server restriction itself (far more expensively, both for
lack of caching and because it would be repeated for each combination of
input sub-joins), or else risk nasty bugs. Anyone who's really intent on
doing something outside this restriction can always use the
set_join_pathlist_hook.
* Rename fdw_ps_tlist/custom_ps_tlist to fdw_scan_tlist/custom_scan_tlist
to better reflect what they're for, and allow these custom scan tlists
to be used even for base relations.
* Change make_foreignscan() API to include passing the fdw_scan_tlist
value, since the FDW is required to set that. Backwards compatibility
doesn't seem like an adequate reason to expect FDWs to set it in some
ad-hoc extra step, and anyway existing FDWs can just pass NIL.
* Change the API of path-generating subroutines of add_paths_to_joinrel,
and in particular that of GetForeignJoinPaths and set_join_pathlist_hook,
so that various less-used parameters are passed in a struct rather than
as separate parameter-list entries. The objective here is to reduce the
probability that future additions to those parameter lists will result in
source-level API breaks for users of these hooks. It's possible that this
is even a small win for the core code, since most CPU architectures can't
pass more than half a dozen parameters efficiently anyway. I kept root,
joinrel, outerrel, innerrel, and jointype as separate parameters to reduce
code churn in joinpath.c --- in particular, putting jointype into the
struct would have been problematic because of the subroutines' habit of
changing their local copies of that variable.
* Avoid ad-hocery in ExecAssignScanProjectionInfo. It was probably all
right for it to know about IndexOnlyScan, but if the list is to grow
we should refactor the knowledge out to the callers.
* Restore nodeForeignscan.c's previous use of the relcache to avoid
extra GetFdwRoutine lookups for base-relation scans.
* Lots of cleanup of documentation and missed comments. Re-order some
code additions into more logical places.
The newly added ON CONFLICT clause allows to specify an alternative to
raising a unique or exclusion constraint violation error when inserting.
ON CONFLICT refers to constraints that can either be specified using a
inference clause (by specifying the columns of a unique constraint) or
by naming a unique or exclusion constraint. DO NOTHING avoids the
constraint violation, without touching the pre-existing row. DO UPDATE
SET ... [WHERE ...] updates the pre-existing tuple, and has access to
both the tuple proposed for insertion and the existing tuple; the
optional WHERE clause can be used to prevent an update from being
executed. The UPDATE SET and WHERE clauses have access to the tuple
proposed for insertion using the "magic" EXCLUDED alias, and to the
pre-existing tuple using the table name or its alias.
This feature is often referred to as upsert.
This is implemented using a new infrastructure called "speculative
insertion". It is an optimistic variant of regular insertion that first
does a pre-check for existing tuples and then attempts an insert. If a
violating tuple was inserted concurrently, the speculatively inserted
tuple is deleted and a new attempt is made. If the pre-check finds a
matching tuple the alternative DO NOTHING or DO UPDATE action is taken.
If the insertion succeeds without detecting a conflict, the tuple is
deemed inserted.
To handle the possible ambiguity between the excluded alias and a table
named excluded, and for convenience with long relation names, INSERT
INTO now can alias its target table.
Bumps catversion as stored rules change.
Author: Peter Geoghegan, with significant contributions from Heikki
Linnakangas and Andres Freund. Testing infrastructure by Jeff Janes.
Reviewed-By: Heikki Linnakangas, Andres Freund, Robert Haas, Simon Riggs,
Dean Rasheed, Stephen Frost and many others.
Previously, relation range table entries used a single Bitmapset field
representing which columns required either UPDATE or INSERT privileges,
despite the fact that INSERT and UPDATE privileges are separately
cataloged, and may be independently held. As statements so far required
either insert or update privileges but never both, that was
sufficient. The required permission could be inferred from the top level
statement run.
The upcoming INSERT ... ON CONFLICT UPDATE feature needs to
independently check for both privileges in one statement though, so that
is not sufficient anymore.
Bumps catversion as stored rules change.
Author: Peter Geoghegan
Reviewed-By: Andres Freund
This makes the executor code more consistent. It also removes
an apparently superfluous NULL test in nodeGroup.c.
Qingqing Zhou, reviewed by Tom Lane, and further revised by me.
Foreign data wrappers can use this capability for so-called "join
pushdown"; that is, instead of executing two separate foreign scans
and then joining the results locally, they can generate a path which
performs the join on the remote server and then is scanned locally.
This commit does not extend postgres_fdw to take advantage of this
capability; it just provides the infrastructure.
Custom scan providers can use this in a similar way. Previously,
it was only possible for a custom scan provider to scan a single
relation. Now, it can scan an entire join tree, provided of course
that it knows how to produce the same results that the join would
have produced if executed normally.
KaiGai Kohei, reviewed by Shigeru Hanada, Ashutosh Bapat, and me.
This does four basic things. First, it provides convenience routines
to coordinate the startup and shutdown of parallel workers. Second,
it synchronizes various pieces of state (e.g. GUCs, combo CID
mappings, transaction snapshot) from the parallel group leader to the
worker processes. Third, it prohibits various operations that would
result in unsafe changes to that state while parallelism is active.
Finally, it propagates events that would result in an ErrorResponse,
NoticeResponse, or NotifyResponse message being sent to the client
from the parallel workers back to the master, from which they can then
be sent on to the client.
Robert Haas, Amit Kapila, Noah Misch, Rushabh Lathia, Jeevan Chalke.
Suggestions and review from Andres Freund, Heikki Linnakangas, Noah
Misch, Simon Riggs, Euler Taveira, and Jim Nasby.
The RLS capability is built on top of the WITH CHECK OPTION
system which was added for auto-updatable views, however, unlike
WCOs on views (which are mandated by the SQL spec to not fire until
after all other constraints and checks are done), it makes much more
sense for RLS checks to happen earlier than constraint and uniqueness
checks.
This patch reworks the structure which holds the WCOs a bit to be
explicitly either VIEW or RLS checks and the RLS-related checks are
done prior to the constraint and uniqueness checks. This also allows
better error reporting as we are now reporting when a violation is due
to a WITH CHECK OPTION and when it's due to an RLS policy violation,
which was independently noted by Craig Ringer as being confusing.
The documentation is also updated to include a paragraph about when RLS
WITH CHECK handling is performed, as there have been a number of
questions regarding that and the documentation was previously silent on
the matter.
Author: Dean Rasheed, with some kabitzing and comment changes by me.
ExecOpenScanRelation assumed that any relation listed in the ExecRowMark
list has been locked by InitPlan; but this is not true if the rel's
markType is ROW_MARK_COPY, which is possible if it's a foreign table.
In most (possibly all) cases, failure to acquire a lock here isn't really
problematic because the parser, planner, or plancache would have taken the
appropriate lock already. In principle though it might leave us vulnerable
to working with a relation that we hold no lock on, and in any case if the
executor isn't depending on previously-taken locks otherwise then it should
not do so for ROW_MARK_COPY relations.
Noted by Etsuro Fujita. Back-patch to all active versions, since the
inconsistency has been there a long time. (It's almost certainly
irrelevant in 9.0, since that predates foreign tables, but the code's
still wrong on its own terms.)
Foreign tables can now be inheritance children, or parents. Much of the
system was already ready for this, but we had to fix a few things of
course, mostly in the area of planner and executor handling of row locks.
As side effects of this, allow foreign tables to have NOT VALID CHECK
constraints (and hence to accept ALTER ... VALIDATE CONSTRAINT), and to
accept ALTER SET STORAGE and ALTER SET WITH/WITHOUT OIDS. Continuing to
disallow these things would've required bizarre and inconsistent special
cases in inheritance behavior. Since foreign tables don't enforce CHECK
constraints anyway, a NOT VALID one is a complete no-op, but that doesn't
mean we shouldn't allow it. And it's possible that some FDWs might have
use for SET STORAGE or SET WITH OIDS, though doubtless they will be no-ops
for most.
An additional change in support of this is that when a ModifyTable node
has multiple target tables, they will all now be explicitly identified
in EXPLAIN output, for example:
Update on pt1 (cost=0.00..321.05 rows=3541 width=46)
Update on pt1
Foreign Update on ft1
Foreign Update on ft2
Update on child3
-> Seq Scan on pt1 (cost=0.00..0.00 rows=1 width=46)
-> Foreign Scan on ft1 (cost=100.00..148.03 rows=1170 width=46)
-> Foreign Scan on ft2 (cost=100.00..148.03 rows=1170 width=46)
-> Seq Scan on child3 (cost=0.00..25.00 rows=1200 width=46)
This was done mainly to provide an unambiguous place to attach "Remote SQL"
fields, but it is useful for inherited updates even when no foreign tables
are involved.
Shigeru Hanada and Etsuro Fujita, reviewed by Ashutosh Bapat and Kyotaro
Horiguchi, some additional hacking by me
The ROW_MARK_COPY path in EvalPlanQualFetchRowMarks() was just setting
tableoid to InvalidOid, I think on the assumption that the referenced
RTE must be a subquery or other case without a meaningful OID. However,
foreign tables also use this code path, and they do have meaningful
table OIDs; so failure to set the tuple field can lead to user-visible
misbehavior. Fix that by fetching the appropriate OID from the range
table.
There's still an issue about whether CTID can ever have a meaningful
value in this case; at least with postgres_fdw foreign tables, it does.
But that is a different problem that seems to require a significantly
different patch --- it's debatable whether postgres_fdw really wants to
use this code path at all.
Simplified version of a patch by Etsuro Fujita, who also noted the
problem to begin with. The issue can be demonstrated in all versions
having FDWs, so back-patch to 9.1.
Previously, we cached domain constraints for the life of a query, or
really for the life of the FmgrInfo struct that was used to invoke
domain_in() or domain_check(). But plpgsql (and probably other places)
are set up to cache such FmgrInfos for the whole lifespan of a session,
which meant they could be enforcing really stale sets of constraints.
On the other hand, searching pg_constraint once per query gets kind of
expensive too: testing says that as much as half the runtime of a
trivial query such as "SELECT 0::domaintype" went into that.
To fix this, delegate the responsibility for tracking a domain's
constraints to the typcache, which has the infrastructure needed to
detect syscache invalidation events that signal possible changes.
This not only removes unnecessary repeat reads of pg_constraint,
but ensures that we never apply stale constraint data: whatever we
use is the current data according to syscache rules.
Unfortunately, the current configuration of the system catalogs means
we have to flush cached domain-constraint data whenever either pg_type
or pg_constraint changes, which happens rather a lot (eg, creation or
deletion of a temp table will do it). It might be worth rearranging
things to split pg_constraint into two catalogs, of which the domain
constraint one would probably be very low-traffic. That's a job for
another patch though, and in any case this patch should improve matters
materially even with that handicap.
This patch makes use of the recently-added memory context reset callback
feature to manage the lifespan of domain constraint caches, so that we
don't risk deleting a cache that might be in the midst of evaluation.
Although this is a bug fix as well as a performance improvement, no
back-patch. There haven't been many if any field complaints about
stale domain constraint checks, so it doesn't seem worth taking the
risk of modifying data structures as basic as MemoryContexts in back
branches.
Previously, each new array created a new memory context that started
out at 8kB. This is incredibly wasteful when there are lots of small
groups of just a few elements each.
Change initArrayResult() and friends to accept a "subcontext" argument
to indicate whether the caller wants the ArrayBuildState allocated in
a new subcontext or not. If not, it can no longer be released
separately from the rest of the memory context.
Fixes bug report by Frank van Vugt on 2013-10-19.
Tomas Vondra. Reviewed by Ali Akbar, Tom Lane, and me.
This requires changing quite a few places that were depending on
sizeof(HeapTupleHeaderData), but it seems for the best.
Michael Paquier, some adjustments by me
Replace some bogus "x[1]" declarations with "x[FLEXIBLE_ARRAY_MEMBER]".
Aside from being more self-documenting, this should help prevent bogus
warnings from static code analyzers and perhaps compiler misoptimizations.
This patch is just a down payment on eliminating the whole problem, but
it gets rid of a lot of easy-to-fix cases.
Note that the main problem with doing this is that one must no longer rely
on computing sizeof(the containing struct), since the result would be
compiler-dependent. Instead use offsetof(struct, lastfield). Autoconf
also warns against spelling that offsetof(struct, lastfield[0]).
Michael Paquier, review and additional fixes by me.
The four functions array_ref, array_set, array_get_slice, array_set_slice
have traditionally declared their array inputs and results as being of type
"ArrayType *". This is a lie, and has been since Berkeley days, because
they actually also support "fixed-length array" types such as "name" and
"point"; not to mention that the inputs could be toasted. These values
should be declared Datum instead to avoid confusion. The current coding
already risks possible misoptimization by compilers, and it'll get worse
when "expanded" array representations become a valid alternative.
However, there's a fair amount of code using array_ref and array_set with
arrays that *are* known to be ArrayType structures, and there might be more
such places in third-party code. Rather than cluttering those call sites
with PointerGetDatum/DatumGetArrayTypeP cruft, what I did was to rename the
existing functions to array_get_element/array_set_element, fix their
signatures, then reincarnate array_ref/array_set as backwards compatibility
wrappers.
array_get_slice/array_set_slice have no such constituency in the core code,
and probably not in third-party code either, so I just changed their APIs.
In commit bf7ca15875 I introduced an
assumption that an RTE referenced by a whole-row Var must have a valid eref
field. This is false for RTEs constructed by DoCopy, and there are other
places taking similar shortcuts. Perhaps we should make all those places
go through addRangeTableEntryForRelation or its siblings instead of having
ad-hoc logic, but the most reliable fix seems to be to make the new code in
ExecEvalWholeRowVar cope if there's no eref. We can reasonably assume that
there's no need to insert column aliases if no aliases were provided.
Add a regression test case covering this, and also verifying that a sane
column name is in fact available in this situation.
Although the known case only crashes in 9.4 and HEAD, it seems prudent to
back-patch the code change to 9.2, since all the ingredients for a similar
failure exist in the variant patch applied to 9.3 and 9.2.
Per report from Jean-Pierre Pelletier.
If an insertion or update had to wait for another transaction to finish,
because there was another insertion with conflicting key in progress,
we would pass a just-free'd item pointer to XactLockTableWait().
All calls to XactLockTableWait() and MultiXactIdWait() had similar issues.
Some passed a pointer to a buffer in the buffer cache, after already
releasing the lock. The call in EvalPlanQualFetch had already released the
pin too. All but the call in execUtils.c would merely lead to reporting a
bogus ctid, however (or an assertion failure, if enabled).
All the callers that passed HeapTuple->t_data->t_ctid were slightly bogus
anyway: if the tuple was updated (again) in the same transaction, its ctid
field would point to the next tuple in the chain, not the tuple itself.
Backpatch to 9.4, where the 'ctid' argument to XactLockTableWait was added
(in commit f88d4cfc)
While building error messages to return to the user,
BuildIndexValueDescription, ExecBuildSlotValueDescription and
ri_ReportViolation would happily include the entire key or entire row in
the result returned to the user, even if the user didn't have access to
view all of the columns being included.
Instead, include only those columns which the user is providing or which
the user has select rights on. If the user does not have any rights
to view the table or any of the columns involved then no detail is
provided and a NULL value is returned from BuildIndexValueDescription
and ExecBuildSlotValueDescription. Note that, for key cases, the user
must have access to all of the columns for the key to be shown; a
partial key will not be returned.
Further, in master only, do not return any data for cases where row
security is enabled on the relation and row security should be applied
for the user. This required a bit of refactoring and moving of things
around related to RLS- note the addition of utils/misc/rls.c.
Back-patch all the way, as column-level privileges are now in all
supported versions.
This has been assigned CVE-2014-8161, but since the issue and the patch
have already been publicized on pgsql-hackers, there's no point in trying
to hide this commit.
This commit extends the SortSupport infrastructure to allow operator
classes the option to provide abbreviated representations of Datums;
in the case of text, we abbreviate by taking the first few characters
of the strxfrm() blob. If the abbreviated comparison is insufficent
to resolve the comparison, we fall back on the normal comparator.
This can be much faster than the old way of doing sorting if the
first few bytes of the string are usually sufficient to resolve the
comparison.
There is the potential for a performance regression if all of the
strings to be sorted are identical for the first 8+ characters and
differ only in later positions; therefore, the SortSupport machinery
now provides an infrastructure to abort the use of abbreviation if
it appears that abbreviation is producing comparatively few distinct
keys. HyperLogLog, a streaming cardinality estimator, is included in
this commit and used to make that determination for text.
Peter Geoghegan, reviewed by me.
Up to now, the "child" executor state trees generated for EvalPlanQual
rechecks have simply shared the ResultRelInfo arrays used for the original
execution tree. However, this leads to dangling-pointer problems, because
ExecInitModifyTable() is all too willing to scribble on some fields of the
ResultRelInfo(s) even when it's being run in one of those child trees.
This trashes those fields from the perspective of the parent tree, because
even if the generated subtree is logically identical to what was in use in
the parent, it's in a memory context that will go away when we're done
with the child state tree.
We do however want to share information in the direction from the parent
down to the children; in particular, fields such as es_instrument *must*
be shared or we'll lose the stats arising from execution of the children.
So the simplest fix is to make a copy of the parent's ResultRelInfo array,
but not copy any fields back at end of child execution.
Per report from Manuel Kniep. The added isolation test is based on his
example. In an unpatched memory-clobber-enabled build it will reliably
fail with "ctid is NULL" errors in all branches back to 9.1, as a
consequence of junkfilter->jf_junkAttNo being overwritten with $7f7f.
This test cannot be run as-is before that for lack of WITH syntax; but
I have no doubt that some variant of this problem can arise in older
branches, so apply the code change all the way back.
In READ COMMITTED mode, if a SELECT FOR UPDATE discovers it has to redo
WHERE-clause checking on rows that have been updated since the SELECT's
snapshot, it invokes EvalPlanQual processing to do that. If this first
occurs within a non-first child table of an inheritance tree, the previous
coding could accidentally re-return a matching row from an earlier,
already-scanned child table. (And, to add insult to injury, I think this
could make it miss returning a row that should have been returned, if the
updated row that this happens on should still have passed the WHERE qual.)
Per report from Kyotaro Horiguchi; the added isolation test is based on his
test case.
This has been broken for quite awhile, so back-patch to all supported
branches.
This patch adds a way of iterating through the members of a bitmapset
nondestructively, unlike the old way with bms_first_member(). While
bms_next_member() is very slightly slower than bms_first_member()
(at least for typical-size bitmapsets), eliminating the need to palloc
and pfree a temporary copy of the target bitmapset is a significant win.
So this method should be preferred in all cases where a temporary copy
would be necessary.
Tom Lane, with suggestions from Dean Rasheed and David Rowley
As pointed out by Robert, we should really have named pg_rowsecurity
pg_policy, as the objects stored in that catalog are policies. This
patch fixes that and updates the column names to start with 'pol' to
match the new catalog name.
The security consideration for COPY with row level security, also
pointed out by Robert, has also been addressed by remembering and
re-checking the OID of the relation initially referenced during COPY
processing, to make sure it hasn't changed under us by the time we
finish planning out the query which has been built.
Robert and Alvaro also commented on missing OCLASS and OBJECT entries
for POLICY (formerly ROWSECURITY or POLICY, depending) in various
places. This patch fixes that too, which also happens to add the
ability to COMMENT on policies.
In passing, attempt to improve the consistency of messages, comments,
and documentation as well. This removes various incarnations of
'row-security', 'row-level security', 'Row-security', etc, in favor
of 'policy', 'row level security' or 'row_security' as appropriate.
Happy Thanksgiving!
These cases formerly failed with errors about "could not find array type
for data type". Now they yield arrays of the same element type and one
higher dimension.
The implementation involves creating functions with API similar to the
existing accumArrayResult() family. I (tgl) also extended the base family
by adding an initArrayResult() function, which allows callers to avoid
special-casing the zero-inputs case if they just want an empty array as
result. (Not all do, so the previous calling convention remains valid.)
This allowed simplifying some existing code in xml.c and plperl.c.
Ali Akbar, reviewed by Pavel Stehule, significantly modified by me
There seems no prospect that any of this will ever be useful, and indeed
it's questionable whether some of it would work if it ever got called;
it's certainly not been exercised in a very long time, if ever. So let's
get rid of it, and make the comments about mark/restore in execAmi.c less
wishy-washy.
The mark/restore support for Result nodes is also currently dead code,
but that's due to planner limitations not because it's impossible that
it could be useful. So I left it in.
Get rid of the pernicious entanglement between planner and executor headers
introduced by commit 0b03e5951b.
Also, rearrange the CustomFoo struct/typedef definitions so that all the
typedef names are seen as used by the compiler. Without this pgindent
will mess things up a bit, which is not so important perhaps, but it also
removes a bizarre discrepancy between the declaration arrangement used for
CustomExecMethods and that used for CustomScanMethods and
CustomPathMethods.
Clean up the commentary around ExecSupportsMarkRestore to reflect the
rather large change in its API.
Const-ify register_custom_path_provider's argument. This necessitates
casting away const in the function, but that seems better than forcing
callers of the function to do so (or else not const-ify their method
pointer structs, which was sort of the whole point).
De-export fix_expr_common. I don't like the exporting of fix_scan_expr
or replace_nestloop_params either, but this one surely has got little
excuse.
execCurrent.c's search_plan_tree() must recognize a CustomScan on the
target relation. This would only be helpful for custom providers that
support CurrentOfExpr quals, which is probably a bit far-fetched, but
it's not impossible I think. But even without assuming that, we need
to recognize a scanned-relation match so that we will properly throw
error if the desired relation is being scanned with both a CustomScan
and a regular scan (ie, self-join).
Also recognize ForeignScanState for similar reasons. Supporting WHERE
CURRENT OF on a foreign table is probably even more far-fetched than
it is for custom scans, but I think in principle you could do it with
postgres_fdw (or another FDW that supports the ctid column). This
would be a back-patchable bug fix if existing FDWs handled CurrentOfExpr,
but I doubt any do so I won't bother back-patching.
This only happens if a client issues a Parse message with an empty query
string, which is a bit odd; but since it is explicitly called out as legal
by our FE/BE protocol spec, we'd probably better continue to allow it.
Fix by adding tests everywhere that the raw_parse_tree field is passed to
functions that don't or shouldn't accept NULL. Also make it clear in the
relevant comments that NULL is an expected case.
This reverts commits a73c9dbab0 and
2e9650cbcf, which fixed specific crash
symptoms by hacking things at what now seems to be the wrong end, ie the
callee functions. Making the callees allow NULL is superficially more
robust, but it's not always true that there is a defensible thing for the
callee to do in such cases. The caller has more context and is better
able to decide what the empty-query case ought to do.
Per followup discussion of bug #11335. Back-patch to 9.2. The code
before that is sufficiently different that it would require development
of a separate patch, which doesn't seem worthwhile for what is believed
to be an essentially cosmetic change.
At one time it wasn't terribly important what column names were associated
with the fields of a composite Datum, but since the introduction of
operations like row_to_json(), it's important that looking up the rowtype
ID embedded in the Datum returns the column names that users would expect.
That did not work terribly well before this patch: you could get the column
names of the underlying table, or column aliases from any level of the
query, depending on minor details of the plan tree. You could even get
totally empty field names, which is disastrous for cases like row_to_json().
To fix this for whole-row Vars, look to the RTE referenced by the Var, and
make sure its column aliases are applied to the rowtype associated with
the result Datums. This is a tad scary because we might have to return
a transient RECORD type even though the Var is declared as having some
named rowtype. In principle it should be all right because the record
type will still be physically compatible with the named rowtype; but
I had to weaken one Assert in ExecEvalConvertRowtype, and there might be
third-party code containing similar assumptions.
Similarly, RowExprs have to be willing to override the column names coming
from a named composite result type and produce a RECORD when the column
aliases visible at the site of the RowExpr differ from the underlying
table's column names.
In passing, revert the decision made in commit 398f70ec07 to add
an alias-list argument to ExecTypeFromExprList: better to provide that
functionality in a separate function. This also reverts most of the code
changes in d685814835, which we don't need because we're no longer
depending on the tupdesc found in the child plan node's result slot to be
blessed.
Back-patch to 9.4, but not earlier, since this solution changes the results
in some cases that users might not have realized were buggy. We'll apply a
more restricted form of this patch in older branches.
This allows extension modules to define their own methods for
scanning a relation, and get the core code to use them. It's
unclear as yet how much use this capability will find, but we
won't find out if we never commit it.
KaiGai Kohei, reviewed at various times and in various levels
of detail by Shigeru Hanada, Tom Lane, Andres Freund, Álvaro
Herrera, and myself.
If we expect batching at the very beginning, we size nbuckets for
"full work_mem" (see how many tuples we can get into work_mem,
while not breaking NTUP_PER_BUCKET threshold).
If we expect to be fine without batching, we start with the 'right'
nbuckets and track the optimal nbuckets as we go (without actually
resizing the hash table). Once we hit work_mem (considering the
optimal nbuckets value), we keep the value.
At the end of the first batch, we check whether (nbuckets !=
nbuckets_optimal) and resize the hash table if needed. Also, we
keep this value for all batches (it's OK because it assumes full
work_mem, and it makes the batchno evaluation trivial). So the
resize happens only once.
There could be cases where it would improve performance to allow
the NTUP_PER_BUCKET threshold to be exceeded to keep everything in
one batch rather than spilling to a second batch, but attempts to
generate such a case have so far been unsuccessful; that issue may
be addressed with a follow-on patch after further investigation.
Tomas Vondra with minor format and comment cleanup by me
Reviewed by Robert Haas, Heikki Linnakangas, and Kevin Grittner
This clause changes the behavior of SELECT locking clauses in the
presence of locked rows: instead of causing a process to block waiting
for the locks held by other processes (or raise an error, with NOWAIT),
SKIP LOCKED makes the new reader skip over such rows. While this is not
appropriate behavior for general purposes, there are some cases in which
it is useful, such as queue-like tables.
Catalog version bumped because this patch changes the representation of
stored rules.
Reviewed by Craig Ringer (based on a previous attempt at an
implementation by Simon Riggs, who also provided input on the syntax
used in the current patch), David Rowley, and Álvaro Herrera.
Author: Thomas Munro
Building on the updatable security-barrier views work, add the
ability to define policies on tables to limit the set of rows
which are returned from a query and which are allowed to be added
to a table. Expressions defined by the policy for filtering are
added to the security barrier quals of the query, while expressions
defined to check records being added to a table are added to the
with-check options of the query.
New top-level commands are CREATE/ALTER/DROP POLICY and are
controlled by the table owner. Row Security is able to be enabled
and disabled by the owner on a per-table basis using
ALTER TABLE .. ENABLE/DISABLE ROW SECURITY.
Per discussion, ROW SECURITY is disabled on tables by default and
must be enabled for policies on the table to be used. If no
policies exist on a table with ROW SECURITY enabled, a default-deny
policy is used and no records will be visible.
By default, row security is applied at all times except for the
table owner and the superuser. A new GUC, row_security, is added
which can be set to ON, OFF, or FORCE. When set to FORCE, row
security will be applied even for the table owner and superusers.
When set to OFF, row security will be disabled when allowed and an
error will be thrown if the user does not have rights to bypass row
security.
Per discussion, pg_dump sets row_security = OFF by default to ensure
that exports and backups will have all data in the table or will
error if there are insufficient privileges to bypass row security.
A new option has been added to pg_dump, --enable-row-security, to
ask pg_dump to export with row security enabled.
A new role capability, BYPASSRLS, which can only be set by the
superuser, is added to allow other users to be able to bypass row
security using row_security = OFF.
Many thanks to the various individuals who have helped with the
design, particularly Robert Haas for his feedback.
Authors include Craig Ringer, KaiGai Kohei, Adam Brightwell, Dean
Rasheed, with additional changes and rework by me.
Reviewers have included all of the above, Greg Smith,
Jeff McCormick, and Robert Haas.
Since this makes the bucket headers use ~10x as much memory, properly
account for that memory when we figure out whether everything fits
in work_mem. This might result in some cases that previously used
only a single batch getting split into multiple batches, but it's
unclear as yet whether we need defenses against that case, and if so,
what the shape of those defenses should be.
It's worth noting that even in these edge cases, users should still be
no worse off than they would have been last week, because commit
45f6240a8f saved a big pile of memory
on exactly the same workloads.
Tomas Vondra, reviewed and somewhat revised by me.
Instead of palloc'ing each HashJoinTuple individually, allocate 32kB chunks
and pack the tuples densely in the chunks. This avoids the AllocChunk
header overhead, and the space wasted by standard allocator's habit of
rounding sizes up to the nearest power of two.
This doesn't contain any planner changes, because the planner's estimate of
memory usage ignores the palloc overhead. Now that the overhead is smaller,
the planner's estimates are in fact more accurate.
Tomas Vondra, reviewed by Robert Haas.
If SELECT FOR UPDATE NOWAIT tries to lock a tuple that is concurrently
being updated, it might fail to honor its NOWAIT specification and block
instead of raising an error.
Fix by adding a no-wait flag to EvalPlanQualFetch which it can pass down
to heap_lock_tuple; also use it in EvalPlanQualFetch itself to avoid
blocking while waiting for a concurrent transaction.
Authors: Craig Ringer and Thomas Munro, tweaked by Álvaro
http://www.postgresql.org/message-id/51FB6703.9090801@2ndquadrant.com
Per Thomas Munro in the course of his SKIP LOCKED feature submission,
who also provided one of the isolation test specs.
Backpatch to 9.4, because that's as far back as it applies without
conflicts (although the bug goes all the way back). To that branch also
backpatch Thomas Munro's new NOWAIT test cases, committed in master by
Heikki as commit 9ee16b49f0 .
This could be useful for datatypes like text, where we might want
to optimize for some collations but not others. However, this patch
doesn't introduce any new sortsupport functions that work this way;
it merely revises the code so that future patches may do so.
Patch by me. Review by Peter Geoghegan.
ExecEvalWholeRowVar incorrectly supposed that it could "bless" the source
TupleTableSlot just once per query. But if the input is coming from an
Append (or, perhaps, other cases?) more than one slot might be returned
over the query run. This led to "record type has not been registered"
errors when a composite datum was extracted from a non-blessed slot.
This bug has been there a long time; I guess it escaped notice because when
dealing with subqueries the planner tends to expand whole-row Vars into
RowExprs, which don't have the same problem. It is possible to trigger
the problem in all active branches, though, as illustrated by the added
regression test.
The previous design exposed the input and output ExprContexts of the
Agg plan node, but work on grouping sets has suggested that we'll regret
doing that. Instead provide more narrowly-defined APIs that can be
implemented in multiple ways, namely a way to get a short-term memory
context and a way to register an aggregate shutdown callback.
Back-patch to 9.4 where the bad APIs were introduced, since we don't
want third-party code using these APIs and then having to change in 9.5.
Andrew Gierth
ExecMakeTableFunctionResult evaluated the arguments for a function-in-FROM
in the query-lifespan memory context. This is insignificant in simple
cases where the function relation is scanned only once; but if the function
is in a sub-SELECT or is on the inside of a nested loop, any memory
consumed during argument evaluation can add up quickly. (The potential for
trouble here had been foreseen long ago, per existing comments; but we'd
not previously seen a complaint from the field about it.) To fix, create
an additional temporary context just for this purpose.
Per an example from MauMau. Back-patch to all active branches.
This SQL-standard feature allows a sub-SELECT yielding multiple columns
(but only one row) to be used to compute the new values of several columns
to be updated. While the same results can be had with an independent
sub-SELECT per column, such a workaround can require a great deal of
duplicated computation.
The standard actually says that the source for a multi-column assignment
could be any row-valued expression. The implementation used here is
tightly tied to our existing sub-SELECT support and can't handle other
cases; the Bison grammar would have some issues with them too. However,
I don't feel too bad about this since other cases can be converted into
sub-SELECTs. For instance, "SET (a,b,c) = row_valued_function(x)" could
be written "SET (a,b,c) = (SELECT * FROM row_valued_function(x))".
Index-only scans avoid taking a lock on the VM buffer, which would
cause a lot of contention. To be correct, that requires some intricate
assumptions that weren't completely documented in the previous
comment.
Reviewed by Robert Haas.
If we have an array of records stored on disk, the individual record fields
cannot contain out-of-line TOAST pointers: the tuptoaster.c mechanisms are
only prepared to deal with TOAST pointers appearing in top-level fields of
a stored row. The same applies for ranges over composite types, nested
composites, etc. However, the existing code only took care of expanding
sub-field TOAST pointers for the case of nested composites, not for other
structured types containing composites. For example, given a command such
as
UPDATE tab SET arraycol = ARRAY[(ROW(x,42)::mycompositetype] ...
where x is a direct reference to a field of an on-disk tuple, if that field
is long enough to be toasted out-of-line then the TOAST pointer would be
inserted as-is into the array column. If the source record for x is later
deleted, the array field value would become a dangling pointer, leading
to errors along the line of "missing chunk number 0 for toast value ..."
when the value is referenced. A reproducible test case for this was
provided by Jan Pecek, but it seems likely that some of the "missing chunk
number" reports we've heard in the past were caused by similar issues.
Code-wise, the problem is that PG_DETOAST_DATUM() is not adequate to
produce a self-contained Datum value if the Datum is of composite type.
Seen in this light, the problem is not just confined to arrays and ranges,
but could also affect some other places where detoasting is done in that
way, for example form_index_tuple().
I tried teaching the array code to apply toast_flatten_tuple_attribute()
along with PG_DETOAST_DATUM() when the array element type is composite,
but this was messy and imposed extra cache lookup costs whether or not any
TOAST pointers were present, indeed sometimes when the array element type
isn't even composite (since sometimes it takes a typcache lookup to find
that out). The idea of extending that approach to all the places that
currently use PG_DETOAST_DATUM() wasn't attractive at all.
This patch instead solves the problem by decreeing that composite Datum
values must not contain any out-of-line TOAST pointers in the first place;
that is, we expand out-of-line fields at the point of constructing a
composite Datum, not at the point where we're about to insert it into a
larger tuple. This rule is applied only to true composite Datums, not
to tuples that are being passed around the system as tuples, so it's not
as invasive as it might sound at first. With this approach, the amount
of code that has to be touched for a full solution is greatly reduced,
and added cache lookup costs are avoided except when there actually is
a TOAST pointer that needs to be inlined.
The main drawback of this approach is that we might sometimes dereference
a TOAST pointer that will never actually be used by the query, imposing a
rather large cost that wasn't there before. On the other side of the coin,
if the field value is used multiple times then we'll come out ahead by
avoiding repeat detoastings. Experimentation suggests that common SQL
coding patterns are unaffected either way, though. Applications that are
very negatively affected could be advised to modify their code to not fetch
columns they won't be using.
In future, we might consider reverting this solution in favor of detoasting
only at the point where data is about to be stored to disk, using some
method that can drill down into multiple levels of nested structured types.
That will require defining new APIs for structured types, though, so it
doesn't seem feasible as a back-patchable fix.
Note that this patch changes HeapTupleGetDatum() from a macro to a function
call; this means that any third-party code using that macro will not get
protection against creating TOAST-pointer-containing Datums until it's
recompiled. The same applies to any uses of PG_RETURN_HEAPTUPLEHEADER().
It seems likely that this is not a big problem in practice: most of the
tuple-returning functions in core and contrib produce outputs that could
not possibly be toasted anyway, and the same probably holds for third-party
extensions.
This bug has existed since TOAST was invented, so back-patch to all
supported branches.
Before 9.4, such an aggregate couldn't be declared, because its final
function would have to have polymorphic result type but no polymorphic
argument, which CREATE FUNCTION would quite properly reject. The
ordered-set-aggregate patch found a workaround: allow the final function
to be declared as accepting additional dummy arguments that have types
matching the aggregate's regular input arguments. However, we failed
to notice that this problem applies just as much to regular aggregates,
despite the fact that we had a built-in regular aggregate array_agg()
that was known to be undeclarable in SQL because its final function
had an illegal signature. So what we should have done, and what this
patch does, is to decouple the extra-dummy-arguments behavior from
ordered-set aggregates and make it generally available for all aggregate
declarations. We have to put this into 9.4 rather than waiting till
later because it slightly alters the rules for declaring ordered-set
aggregates.
The patch turned out a bit bigger than I'd hoped because it proved
necessary to record the extra-arguments option in a new pg_aggregate
column. I'd thought we could just look at the final function's pronargs
at runtime, but that didn't work well for variadic final functions.
It's probably just as well though, because it simplifies life for pg_dump
to record the option explicitly.
While at it, fix array_agg() to have a valid final-function signature,
and add an opr_sanity test to notice future deviations from polymorphic
consistency. I also marked the percentile_cont() aggregates as not
needing extra arguments, since they don't.
Repositioning the tuplestore seek pointer in window_gettupleslot() turns
out to be a very significant expense when the window frame is sizable and
the frame end can move. To fix, introduce a tuplestore function for
skipping an arbitrary number of tuples in one call, parallel to the one we
introduced for tuplesort objects in commit 8d65da1f. This reduces the cost
of window_gettupleslot() to O(1) if the tuplestore has not spilled to disk.
As in the previous commit, I didn't try to do any real optimization of
tuplestore_skiptuples for the case where the tuplestore has spilled to
disk. There is probably no practical way to get the cost to less than O(N)
anyway, but perhaps someone can think of something later.
Also fix PersistHoldablePortal() to make use of this API now that we have
it.
Based on a suggestion by Dean Rasheed, though this turns out not to look
much like his patch.
Until now, when executing an aggregate function as a window function
within a window with moving frame start (that is, any frame start mode
except UNBOUNDED PRECEDING), we had to recalculate the aggregate from
scratch each time the frame head moved. This patch allows an aggregate
definition to include an alternate "moving aggregate" implementation
that includes an inverse transition function for removing rows from
the aggregate's running state. As long as this can be done successfully,
runtime is proportional to the total number of input rows, rather than
to the number of input rows times the average frame length.
This commit includes the core infrastructure, documentation, and regression
tests using user-defined aggregates. Follow-on commits will update some
of the built-in aggregates to use this feature.
David Rowley and Florian Pflug, reviewed by Dean Rasheed; additional
hacking by me
This covers all the SQL-standard trigger types supported for regular
tables; it does not cover constraint triggers. The approach for
acquiring the old row mirrors that for view INSTEAD OF triggers. For
AFTER ROW triggers, we spool the foreign tuples to a tuplestore.
This changes the FDW API contract; when deciding which columns to
populate in the slot returned from data modification callbacks, writable
FDWs will need to check for AFTER ROW triggers in addition to checking
for a RETURNING clause.
In support of the feature addition, refactor the TriggerFlags bits and
the assembly of old tuples in ModifyTable.
Ronan Dunklau, reviewed by KaiGai Kohei; some additional hacking by me.
With this in place, a session blocking behind another one because of
tuple locks will get a context line mentioning the relation name, tuple
TID, and operation being done on tuple. For example:
LOG: process 11367 still waiting for ShareLock on transaction 717 after 1000.108 ms
DETAIL: Process holding the lock: 11366. Wait queue: 11367.
CONTEXT: while updating tuple (0,2) in relation "foo"
STATEMENT: UPDATE foo SET value = 3;
Most usefully, the new line is displayed by log entries due to
log_lock_waits, although of course it will be printed by any other log
message as well.
Author: Christian Kruse, some tweaks by Álvaro Herrera
Reviewed-by: Amit Kapila, Andres Freund, Tom Lane, Robert Haas
We should allow this so that matviews can be referenced in UPDATE/DELETE
statements in READ COMMITTED isolation level. The requirement for that
is that a re-fetch by TID will see the same row version the query saw
earlier, which is true of matviews, so there's no reason for the
restriction. Per bug #9398.
Michael Paquier, after a suggestion by me
This feature, building on previous commits, allows the write-ahead log
stream to be decoded into a series of logical changes; that is,
inserts, updates, and deletes and the transactions which contain them.
It is capable of handling decoding even across changes to the schema
of the effected tables. The output format is controlled by a
so-called "output plugin"; an example is included. To make use of
this in a real replication system, the output plugin will need to be
modified to produce output in the format appropriate to that system,
and to perform filtering.
Currently, information can be extracted from the logical decoding
system only via SQL; future commits will add the ability to stream
changes via walsender.
Andres Freund, with review and other contributions from many other
people, including Álvaro Herrera, Abhijit Menon-Sen, Peter Gheogegan,
Kevin Grittner, Robert Haas, Heikki Linnakangas, Fujii Masao, Abhijit
Menon-Sen, Michael Paquier, Simon Riggs, Craig Ringer, and Steve
Singer.
Given a composite-type parameter named x, "$1.*" worked fine, but "x.*"
not so much. This has been broken since named parameter references were
added in commit 9bff0780cf, so patch back
to 9.2. Per bug #9085 from Hardy Falk.
In commit c1352052ef, I implemented an
optimization that assumed that a function's argument expressions would
either always return a set (ie multiple rows), or always not. This is
wrong however: we allow CASE expressions in which some arms return a set
of some type and others just return a scalar of that type. There may be
other examples as well. To fix, replace the run-time test of whether an
argument returned a set with a static precheck (expression_returns_set).
This adds a little bit of query startup overhead, but it seems barely
measurable.
Per bug #8228 from David Johnston. This has been broken since 8.0,
so patch all supported branches.
Keep a pre-initialized FunctionCallInfoData in AggStatePerAggData, and
re-use that at each row instead of doing InitFunctionCallInfoData each
time. This saves only half a dozen assignments and maybe some stack
manipulation, and yet that seems to be good for a percent or two of the
overall query run time for simple aggregates such as count(*). The cost
is that the FunctionCallInfoData (which is about a kilobyte, on 64-bit
machines) stays allocated for the duration of the query instead of being
short-lived stack data. But we're already paying an equivalent space cost
for each regular FuncExpr or OpExpr node, so I don't feel bad about paying
it for aggregate functions. The code seems a little cleaner this way too,
since the number of things passed to advance_transition_function decreases.
This patch introduces generic support for ordered-set and hypothetical-set
aggregate functions, as well as implementations of the instances defined in
SQL:2008 (percentile_cont(), percentile_disc(), rank(), dense_rank(),
percent_rank(), cume_dist()). We also added mode() though it is not in the
spec, as well as versions of percentile_cont() and percentile_disc() that
can compute multiple percentile values in one pass over the data.
Unlike the original submission, this patch puts full control of the sorting
process in the hands of the aggregate's support functions. To allow the
support functions to find out how they're supposed to sort, a new API
function AggGetAggref() is added to nodeAgg.c. This allows retrieval of
the aggregate call's Aggref node, which may have other uses beyond the
immediate need. There is also support for ordered-set aggregates to
install cleanup callback functions, so that they can be sure that
infrastructure such as tuplesort objects gets cleaned up.
In passing, make some fixes in the recently-added support for variadic
aggregates, and make some editorial adjustments in the recent FILTER
additions for aggregates. Also, simplify use of IsBinaryCoercible() by
allowing it to succeed whenever the target type is ANY or ANYELEMENT.
It was inconsistent that it dealt with other polymorphic target types
but not these.
Atri Sharma and Andrew Gierth; reviewed by Pavel Stehule and Vik Fearing,
and rather heavily editorialized upon by Tom Lane
This patch adds the ability to write TABLE( function1(), function2(), ...)
as a single FROM-clause entry. The result is the concatenation of the
first row from each function, followed by the second row from each
function, etc; with NULLs inserted if any function produces fewer rows than
others. This is believed to be a much more useful behavior than what
Postgres currently does with multiple SRFs in a SELECT list.
This syntax also provides a reasonable way to combine use of column
definition lists with WITH ORDINALITY: put the column definition list
inside TABLE(), where it's clear that it doesn't control the ordinality
column as well.
Also implement SQL-compliant multiple-argument UNNEST(), by turning
UNNEST(a,b,c) into TABLE(unnest(a), unnest(b), unnest(c)).
The SQL standard specifies TABLE() with only a single function, not
multiple functions, and it seems to require an implicit UNNEST() which is
not what this patch does. There may be something wrong with that reading
of the spec, though, because if it's right then the spec's TABLE() is just
a pointless alternative spelling of UNNEST(). After further review of
that, we might choose to adopt a different syntax for what this patch does,
but in any case this functionality seems clearly worthwhile.
Andrew Gierth, reviewed by Zoltán Böszörményi and Heikki Linnakangas, and
significantly revised by me
ExecBuildSlotValueDescription() printed "null" for each dropped column in
a row being complained of by ExecConstraints(). This has some sanity in
terms of the underlying implementation, but is of course pretty surprising
to users. To fix, we must pass the target relation's descriptor to
ExecBuildSlotValueDescription(), because the slot descriptor it had been
using doesn't get labeled with attisdropped markers.
Per bug #8408 from Maxim Boguk. Back-patch to 9.2 where the feature of
printing row values in NOT NULL and CHECK constraint violation messages
was introduced.
Michael Paquier and Tom Lane
Historically, printtup() has assumed that it could prevent memory leakage
by pfree'ing the string result of each output function and manually
managing detoasting of toasted values. This amounts to assuming that
datatype output functions never leak any memory internally; an assumption
we've already decided to be bogus elsewhere, for example in COPY OUT.
range_out in particular is known to leak multiple kilobytes per call, as
noted in bug #8573 from Godfried Vanluffelen. While we could go in and fix
that leak, it wouldn't be very notationally convenient, and in any case
there have been and undoubtedly will again be other leaks in other output
functions. So what seems like the best solution is to run the output
functions in a temporary memory context that can be reset after each row,
as we're doing in COPY OUT. Some quick experimentation suggests this is
actually a tad faster than the retail pfree's anyway.
This patch fixes all the variants of printtup, except for debugtup()
which is used in standalone mode. It doesn't seem worth worrying
about query-lifespan leaks in standalone mode, and fixing that case
would be a bit tedious since debugtup() doesn't currently have any
startup or shutdown functions.
While at it, remove manual detoast management from several other
output-function call sites that had copied it from printtup(). This
doesn't make a lot of difference right now, but in view of recent
discussions about supporting "non-flattened" Datums, we're going to
want that code gone eventually anyway.
Back-patch to 9.2 where range_out was introduced. We might eventually
decide to back-patch this further, but in the absence of known major
leaks in older output functions, I'll refrain for now.
A subquery reference to a matview should be allowed by CREATE
MATERIALIZED VIEW WITH NO DATA, just like a direct reference is.
Per bug report from Laurent Sartran.
Backpatch to 9.3.
Previously, arbitray system columns could be mentioned in table
constraints, but they were not correctly checked at runtime, because
the values weren't actually set correctly in the tuple. Since it
seems easy enough to initialize the table OID properly, do that,
and continue allowing that column, but disallow the rest unless and
until someone figures out a way to make them work properly.
No back-patch, because this doesn't seem important enough to take the
risk of destabilizing the back branches. In fact, this will pose a
dump-and-reload hazard for those upgrading from previous versions:
constraints that were accepted before but were not correctly enforced
will now either be enforced correctly or not accepted at all. Either
could result in restore failures, but in practice I think very few
users will notice the difference, since the use case is pretty
marginal anyway and few users will be relying on features that have
not historically worked.
Amit Kapila, reviewed by Rushabh Lathia, with doc changes by me.
There's no inherent reason why an aggregate function can't be variadic
(even VARIADIC ANY) if its transition function can handle the case.
Indeed, this patch to add the feature touches none of the planner or
executor, and little of the parser; the main missing stuff was DDL and
pg_dump support.
It is true that variadic aggregates can create the same sort of ambiguity
about parameters versus ORDER BY keys that was complained of when we
(briefly) had both one- and two-argument forms of string_agg(). However,
the policy formed in response to that discussion only said that we'd not
create any built-in aggregates with varying numbers of arguments, not that
we shouldn't allow users to do it. So the logical extension of that is
we can allow users to make variadic aggregates as long as we're wary about
shipping any such in core.
In passing, this patch allows aggregate function arguments to be named, to
the extent of remembering the names in pg_proc and dumping them in pg_dump.
You can't yet call an aggregate using named-parameter notation. That seems
like a likely future extension, but it'll take some work, and it's not what
this patch is really about. Likewise, there's still some work needed to
make window functions handle VARIADIC fully, but I left that for another
day.
initdb forced because of new aggvariadic field in Aggref parse nodes.
Failing to do so can cause queries to return wrong data, error out or crash.
This requires adding a new binaryheap_reset() method to binaryheap.c,
but that probably should have been there anyway.
Per bug #8410 from Terje Elde. Diagnosis and patch by Andres Freund.
plpgsql often just remembers SPI-result tuple tables in local variables,
and has no mechanism for freeing them if an ereport(ERROR) causes an escape
out of the execution function whose local variable it is. In the original
coding, that wasn't a problem because the tuple table would be cleaned up
when the function's SPI context went away during transaction abort.
However, once plpgsql grew the ability to trap exceptions, repeated
trapping of errors within a function could result in significant
intra-function-call memory leakage, as illustrated in bug #8279 from
Chad Wagner.
We could fix this locally in plpgsql with a bunch of PG_TRY/PG_CATCH
coding, but that would be tedious, probably slow, and prone to bugs of
omission; moreover it would do nothing for similar risks elsewhere.
What seems like a better plan is to make SPI itself responsible for
freeing tuple tables at subtransaction abort. This patch attacks the
problem that way, keeping a list of live tuple tables within each SPI
function context. Currently, such freeing is automatic for tuple tables
made within the failed subtransaction. We might later add a SPI call to
mark a tuple table as not to be freed this way, allowing callers to opt
out; but until someone exhibits a clear use-case for such behavior, it
doesn't seem worth bothering.
A very useful side-effect of this change is that SPI_freetuptable() can
now defend itself against bad calls, such as duplicate free requests;
this should make things more robust in many places. (In particular,
this reduces the risks involved if a third-party extension contains
now-redundant SPI_freetuptable() calls in error cleanup code.)
Even though the leakage problem is of long standing, it seems imprudent
to back-patch this into stable branches, since it does represent an API
semantics change for SPI users. We'll patch this in 9.3, but live with
the leakage in older branches.
As far as I can determine, there's no code in the core distribution
that fails to explicitly set the snapshot of a scan or executor
state. If there is any such code, this will probably cause it to
seg fault; friendlier suggestions were discussed on pgsql-hackers,
but there was no consensus that anything more than this was
needed.
This is another step towards the hoped-for complete removal of
SnapshotNow.
Previously, these functions took a HeapTupleHeader, but upcoming
patches for logical replication will introduce new a new snapshot
type under which the tuple's TID will be used to lookup (CMIN, CMAX)
for visibility determination purposes. This makes that information
available. Code churn is minimal since HeapTupleSatisfiesVisibility
took the HeapTuple anyway, and deferenced it before calling the
satisfies function.
Independently of logical replication, this allows t_tableOid and
t_self to be cross-checked via assertions in tqual.c. This seems
like a useful way to make sure that all callers are setting these
values properly, which has been previously put forward as
desirable.
Andres Freund, reviewed by Álvaro Herrera
For simple views which are automatically updatable, this patch allows
the user to specify what level of checking should be done on records
being inserted or updated. For 'LOCAL CHECK', new tuples are validated
against the conditionals of the view they are being inserted into, while
for 'CASCADED CHECK' the new tuples are validated against the
conditionals for all views involved (from the top down).
This option is part of the SQL specification.
Dean Rasheed, reviewed by Pavel Stehule
This is SQL-standard with a few extensions, namely support for
subqueries and outer references in clause expressions.
catversion bump due to change in Aggref and WindowFunc.
David Fetter, reviewed by Dean Rasheed.
This allows reads to continue without any blocking while a REFRESH
runs. The new data appears atomically as part of transaction
commit.
Review questioned the Assert that a matview was not a system
relation. This will be addressed separately.
Reviewed by Hitoshi Harada, Robert Haas, Andres Freund.
Merged after review with security patch f3ab5d4.
SnapshotNow scans have the undesirable property that, in the face of
concurrent updates, the scan can fail to see either the old or the new
versions of the row. In many cases, we work around this by requiring
DDL operations to hold AccessExclusiveLock on the object being
modified; in some cases, the existing locking is inadequate and random
failures occur as a result. This commit doesn't change anything
related to locking, but will hopefully pave the way to allowing lock
strength reductions in the future.
The major issue has held us back from making this change in the past
is that taking an MVCC snapshot is significantly more expensive than
using a static special snapshot such as SnapshotNow. However, testing
of various worst-case scenarios reveals that this problem is not
severe except under fairly extreme workloads. To mitigate those
problems, we avoid retaking the MVCC snapshot for each new scan;
instead, we take a new snapshot only when invalidation messages have
been processed. The catcache machinery already requires that
invalidation messages be sent before releasing the related heavyweight
lock; else other backends might rely on locally-cached data rather
than scanning the catalog at all. Thus, making snapshot reuse
dependent on the same guarantees shouldn't break anything that wasn't
already subtly broken.
Patch by me. Review by Michael Paquier and Andres Freund.
Extend the FDW API (which we already changed for 9.3) so that an FDW can
report whether specific foreign tables are insertable/updatable/deletable.
The default assumption continues to be that they're updatable if the
relevant executor callback function is supplied by the FDW, but finer
granularity is now possible. As a test case, add an "updatable" option to
contrib/postgres_fdw.
This patch also fixes the information_schema views, which previously did
not think that foreign tables were ever updatable, and fixes
view_is_auto_updatable() so that a view on a foreign table can be
auto-updatable.
initdb forced due to changes in information_schema views and the functions
they rely on. This is a bit unfortunate to do post-beta1, but if we don't
change this now then we'll have another API break for FDWs when we do
change it.
Dean Rasheed, somewhat editorialized on by Tom Lane
In a construct like "select plain_function(set_returning_function(...))",
the plain function is applied to each output row of the SRF successively.
If some of the SRF outputs are NULL, and the plain function is strict,
you'd expect to get NULL results for such rows ... but what actually
happened was that such rows were omitted entirely from the result set.
This was due to confusion of this case with what should happen for nested
set-returning functions; a strict SRF is indeed supposed to yield an empty
set for null input. Per bug #8150 from Erwin Brandstetter.
Although this has been broken forever, we're not back-patching because
of the possibility that some apps out there expect the incorrect behavior.
This change should be listed as a possible incompatibility in the 9.3
release notes.
Choose a saner ordering of parameters (adding a new input param after
the output params seemed a bit random), update the function's header
comment to match reality (cmon folks, is this really that hard?),
get rid of useless and sloppily-defined distinction between
PROCESS_UTILITY_SUBCOMMAND and PROCESS_UTILITY_GENERATED.
Move checking for unscannable matviews into ExecOpenScanRelation, which is
a better place for it first because the open relation is already available
(saving a relcache lookup cycle), and second because this eliminates the
problem of telling the difference between rangetable entries that will or
will not be scanned by the query. In particular we can get rid of the
not-terribly-well-thought-out-or-implemented isResultRel field that the
initial matviews patch added to RangeTblEntry.
Also get rid of entirely unnecessary scannability check in the rewriter,
and a bogus decision about whether RefreshMatViewStmt requires a parse-time
snapshot.
catversion bump due to removal of a RangeTblEntry field, which changes
stored rules.
This was due to incomplete implementation of rowcount reporting
for RMV, which was due to initial waffling on whether it should
be provided. It seems unlikely to be a useful or universally
available number as more sophisticated techniques for maintaining
matviews are added, so remove the partial support rather than
completing it.
Per report of Jeevan Chalke, but with a different fix
In most cases, these were just references to the SQL standard in
general. In a few cases, a contrast was made between SQL92 and later
standards -- those have been kept unchanged.
If an FDW fails to take special measures with a CurrentOfExpr, we will
end up trying to execute it as an ordinary qual, which was being treated
as a purely internal failure condition. Provide a more user-oriented
error message for such cases.
The intent was that being populated would, long term, be just one
of the conditions which could affect whether a matview was
scannable; being populated should be necessary but not always
sufficient to scan the relation. Since only CREATE and REFRESH
currently determine the scannability, names and comments
accidentally conflated these concepts, leading to confusion.
Also add missing locking for the SQL function which allows a
test for scannability, and fix a modularity violatiion.
Per complaints from Tom Lane, although its not clear that these
will satisfy his concerns. Hopefully this will at least better
frame the discussion.
This patch adds the core-system infrastructure needed to support updates
on foreign tables, and extends contrib/postgres_fdw to allow updates
against remote Postgres servers. There's still a great deal of room for
improvement in optimization of remote updates, but at least there's basic
functionality there now.
KaiGai Kohei, reviewed by Alexander Korotkov and Laurenz Albe, and rather
heavily revised by Tom Lane.
This saves several catalog lookups per reference. It's not all that
exciting right now, because we'd managed to minimize the number of places
that need to fetch the data; but the upcoming writable-foreign-tables patch
needs this info in a lot more places.
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.
fmgr_sql had been designed on the assumption that the FmgrInfo it's called
with has only query lifespan. This is demonstrably unsafe in connection
with range types, as shown in bug #7881 from Andrew Gierth. Fix things
so that we re-generate the function's cache data if the (sub)transaction
it was made in is no longer active.
Back-patch to 9.2. This might be needed further back, but it's not clear
whether the case can realistically arise without range types, so for now
I'll desist from back-patching further.
Commit af7914c662, which added the TIMING
option to EXPLAIN, had an oversight: if the TIMING option is disabled
then control in InstrStartNode() goes through an elog(DEBUG2) call, which
typically does nothing but takes a noticeable amount of time to do it.
Tweak the logic to avoid that.
In HEAD, also change the elog(DEBUG2)'s in instrument.c to elog(ERROR).
It's not very clear why they weren't like that to begin with, but this
episode shows that not complaining more vociferously about misuse is
likely to do little except allow bugs to remain hidden.
While at it, adjust some code that was making possibly-dangerous
assumptions about flag bits being in the rightmost byte of the
instrument_options word.
Problem reported by Pavel Stehule (via Tomas Vondra).
exec_simple_check_plan and exec_eval_simple_expr attempted to call
GetCachedPlan directly. This meant that if an error was thrown during
planning, the resulting context traceback would not include the line
normally contributed by _SPI_error_callback. This is already inconsistent,
but just to be really odd, a re-execution of the very same expression
*would* show the additional context line, because we'd already have cached
the plan and marked the expression as non-simple.
The problem is easy to demonstrate in 9.2 and HEAD because planning of a
cached plan doesn't occur at all until GetCachedPlan is done. In earlier
versions, it could only be an issue if initial planning had succeeded, then
a replan was forced (already somewhat improbable for a simple expression),
and the replan attempt failed. Since the issue is mainly cosmetic in older
branches anyway, it doesn't seem worth the risk of trying to fix it there.
It is worth fixing in 9.2 since the instability of the context printout can
affect the results of GET STACKED DIAGNOSTICS, as per a recent discussion
on pgsql-novice.
To fix, introduce a SPI function that wraps GetCachedPlan while installing
the correct callback function. Use this instead of calling GetCachedPlan
directly from plpgsql.
Also introduce a wrapper function for extracting a SPI plan's
CachedPlanSource list. This lets us stop including spi_priv.h in
pl_exec.c, which was never a very good idea from a modularity standpoint.
In passing, fix a similar inconsistency that could occur in SPI_cursor_open,
which was also calling GetCachedPlan without setting up a context callback.
This patch addresses the problem that applications currently have to
extract object names from possibly-localized textual error messages,
if they want to know for example which index caused a UNIQUE_VIOLATION
failure. It adds new error message fields to the wire protocol, which
can carry the name of a table, table column, data type, or constraint
associated with the error. (Since the protocol spec has always instructed
clients to ignore unrecognized field types, this should not create any
compatibility problem.)
Support for providing these new fields has been added to just a limited set
of error reports (mainly, those in the "integrity constraint violation"
SQLSTATE class), but we will doubtless add them to more calls in future.
Pavel Stehule, reviewed and extensively revised by Peter Geoghegan, with
additional hacking by Tom Lane.
Since 9.0, the count parameter has only limited the number of tuples
actually returned by the executor. It doesn't affect the behavior of
INSERT/UPDATE/DELETE unless RETURNING is specified, because without
RETURNING, the ModifyTable plan node doesn't return control to execMain.c
for each tuple. And we only check the limit at the top level.
While this behavioral change was unintentional at the time, discussion of
bug #6572 led us to the conclusion that we prefer the new behavior anyway,
and so we should just adjust the docs to match rather than change the code.
Accordingly, do that. Back-patch as far as 9.0 so that the docs match the
code in each branch.
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
SPI_execute() and related functions create a CachedPlan, execute it once,
and immediately discard it, so that the functionality offered by
plancache.c is of no value in this code path. And performance measurements
show that the extra data copying and invalidation checking done by
plancache.c slows down simple queries by 10% or more compared to 9.1.
However, enough of the SPI code is shared with functions that do need plan
caching that it seems impractical to bypass plancache.c altogether.
Instead, let's invent a variant version of cached plans that preserves
99% of the API but doesn't offer any of the actual functionality, nor the
overhead. This puts SPI_execute() performance back on par, or maybe even
slightly better, than it was before. This change should resolve recent
complaints of performance degradation from Dong Ye, Pavel Stehule, and
others.
By avoiding data copying, this change also reduces the amount of memory
needed to execute many-statement SPI_execute() strings, as for instance in
a recent complaint from Tomas Vondra.
An additional benefit of this change is that multi-statement SPI_execute()
query strings are now processed fully serially, that is we complete
execution of earlier statements before running parse analysis and planning
on following ones. This eliminates a long-standing POLA violation, in that
DDL that affects the behavior of a later statement will now behave as
expected.
Back-patch to 9.2, since this was a performance regression compared to 9.1.
(In 9.2, place the added struct fields so as to avoid changing the offsets
of existing fields.)
Heikki Linnakangas and Tom Lane
The dynahash code requires the number of buckets in a hash table to fit
in an int; but since we calculate the desired hash table size dynamically,
there are various scenarios where we might calculate too large a value.
The resulting overflow can lead to infinite loops, division-by-zero
crashes, etc. I (tgl) had previously installed some defenses against that
in commit 299d171652, but that covered only one
call path. Moreover it worked by limiting the request size to work_mem,
but in a 64-bit machine it's possible to set work_mem high enough that the
problem appears anyway. So let's fix the problem at the root by installing
limits in the dynahash.c functions themselves.
Trouble report and patch by Jeff Davis.
This patch makes "simple" views automatically updatable, without the need
to create either INSTEAD OF triggers or INSTEAD rules. "Simple" views
are those classified as updatable according to SQL-92 rules. The rewriter
transforms INSERT/UPDATE/DELETE commands on such views directly into an
equivalent command on the underlying table, which will generally have
noticeably better performance than is possible with either triggers or
user-written rules. A view that has INSTEAD OF triggers or INSTEAD rules
continues to operate the same as before.
For the moment, security_barrier views are not considered simple.
Also, we do not support WITH CHECK OPTION. These features may be
added in future.
Dean Rasheed, reviewed by Amit Kapila
There are probably other places where this can be used, but for now,
this just makes MergeAppend use it, so that this code will have test
coverage. There is other work in the queue that will use this, as
well.
Abhijit Menon-Sen, reviewed by Andres Freund, Robert Haas, Álvaro
Herrera, Tom Lane, and others.
Commit 8cb53654db, which introduced DROP
INDEX CONCURRENTLY, managed to break CREATE INDEX CONCURRENTLY via a poor
choice of catalog state representation. The pg_index state for an index
that's reached the final pre-drop stage was the same as the state for an
index just created by CREATE INDEX CONCURRENTLY. This meant that the
(necessary) change to make RelationGetIndexList ignore about-to-die indexes
also made it ignore freshly-created indexes; which is catastrophic because
the latter do need to be considered in HOT-safety decisions. Failure to
do so leads to incorrect index entries and subsequently wrong results from
queries depending on the concurrently-created index.
To fix, add an additional boolean column "indislive" to pg_index, so that
the freshly-created and about-to-die states can be distinguished. (This
change obviously is only possible in HEAD. This patch will need to be
back-patched, but in 9.2 we'll use a kluge consisting of overloading the
formerly-impossible state of indisvalid = true and indisready = false.)
In addition, change CREATE/DROP INDEX CONCURRENTLY so that the pg_index
flag changes they make without exclusive lock on the index are made via
heap_inplace_update() rather than a normal transactional update. The
latter is not very safe because moving the pg_index tuple could result in
concurrent SnapshotNow scans finding it twice or not at all, thus possibly
resulting in index corruption. This is a pre-existing bug in CREATE INDEX
CONCURRENTLY, which was copied into the DROP code.
In addition, fix various places in the code that ought to check to make
sure that the indexes they are manipulating are valid and/or ready as
appropriate. These represent bugs that have existed since 8.2, since
a failed CREATE INDEX CONCURRENTLY could leave a corrupt or invalid
index behind, and we ought not try to do anything that might fail with
such an index.
Also fix RelationReloadIndexInfo to ensure it copies all the pg_index
columns that are allowed to change after initial creation. Previously we
could have been left with stale values of some fields in an index relcache
entry. It's not clear whether this actually had any user-visible
consequences, but it's at least a bug waiting to happen.
In addition, do some code and docs review for DROP INDEX CONCURRENTLY;
some cosmetic code cleanup but mostly addition and revision of comments.
This will need to be back-patched, but in a noticeably different form,
so I'm committing it to HEAD before working on the back-patch.
Problem reported by Amit Kapila, diagnosis by Pavan Deolassee,
fix by Tom Lane and Andres Freund.
This reverts commit d573e239f0, "Take fewer
snapshots". While that seemed like a good idea at the time, it caused
execution to use a snapshot that had been acquired before locking any of
the tables mentioned in the query. This created user-visible anomalies
that were not present in any prior release of Postgres, as reported by
Tomas Vondra. While this whole area could do with a redesign (since there
are related cases that have anomalies anyway), it doesn't seem likely that
any future patch would be reasonably back-patchable; and we don't want 9.2
to exhibit a behavior that's subtly unlike either past or future releases.
Hence, revert to prior code while we rethink the problem.
This prevents surprising behavior when a FOR EACH ROW trigger
BEFORE UPDATE or BEFORE DELETE directly or indirectly updates or
deletes the the old row. Prior to this patch the requested action
on the row could be silently ignored while all triggered actions
based on the occurence of the requested action could be committed.
One example of how this could happen is if the BEFORE DELETE
trigger for a "parent" row deleted "children" which had trigger
functions to update summary or status data on the parent.
This also prevents similar surprising problems if the query has a
volatile function which updates a target row while it is already
being updated.
There are related issues present in FOR UPDATE cursors and READ
COMMITTED queries which are not handled by this patch. These
issues need further evalution to determine what change, if any, is
needed.
Where the new error messages are generated, in most cases the best
fix will be to move code from the BEFORE trigger to an AFTER
trigger. Where this is not feasible, the trigger can avoid the
error by re-issuing the triggering statement and returning NULL.
Documentation changes will be submitted in a separate patch.
Kevin Grittner and Tom Lane with input from Florian Pflug and
Robert Haas, based on problems encountered during conversion of
Wisconsin Circuit Court trigger logic to plpgsql triggers.
When hashing a subplan like "WHERE (a, b) NOT IN (SELECT x, y FROM ...)",
findPartialMatch() attempted to match rows using the hashtable's internal
equality operators, which of course are for x and y's datatypes. What we
need to use are the potentially cross-type operators for a=x, b=y, etc.
Failure to do that leads to wrong answers or even crashes. The scope for
problems is limited to cases where we have different types with compatible
hash functions (else we'd not be using a hashed subplan), but for example
int4 vs int8 can cause the problem.
Per bug #7597 from Bo Jensen. This has been wrong since the hashed-subplan
code was written, so patch all the way back.
You can now get the number of rows processed by a COPY statement in a
PL/pgSQL function with "GET DIAGNOSTICS x = ROW_COUNT".
Pavel Stehule, reviewed by Amit Kapila, with some editing by me.
Serializable Snapshot Isolation used for serializable transactions
depends on acquiring SIRead locks on all heap relation tuples which
are used to generate the query result, so that a later delete or
update of any of the tuples can flag a read-write conflict between
transactions. This is normally handled in heapam.c, with tuple level
locking. Since an index-only scan avoids heap access in many cases,
building the result from the index tuple, the necessary predicate
locks were not being acquired for all tuples in an index-only scan.
To prevent problems with tuple IDs which are vacuumed and re-used
while the transaction still matters, the xmin of the tuple is part of
the tag for the tuple lock. Since xmin is not available to the
index-only scan for result rows generated from the index tuples, it
is not possible to acquire a tuple-level predicate lock in such
cases, in spite of having the tid. If we went to the heap to get the
xmin value, it would no longer be an index-only scan. Rather than
prohibit index-only scans under serializable transaction isolation,
we acquire an SIRead lock on the page containing the tuple, when it
was not necessary to visit the heap for other reasons.
Backpatch to 9.2.
Kevin Grittner and Tom Lane
This reduces unnecessary exposure of other headers through htup.h, which
is very widely included by many files.
I have chosen to move the function prototypes to the new file as well,
because that means htup.h no longer needs to include tupdesc.h. In
itself this doesn't have much effect in indirect inclusion of tupdesc.h
throughout the tree, because it's also required by execnodes.h; but it's
something to explore in the future, and it seemed best to do the htup.h
change now while I'm busy with it.
Tom Lane
Robert Haas
Bruce Momjian
Teodor Sigaev
Peter Eisentraut
Andres Freund
Stephen Frost
Alvaro Herrera
Kevin Grittner
Heikki Linnakangas
Joe Conway
Simon Riggs
Jeff Davis
Noah Misch
Greg Stark