I failed to account for the possibility that when
ExecAppendAsyncEventWait() notifies multiple async-capable nodes using
postgres_fdw, a preceding node might invoke process_pending_request() to
process a pending asynchronous request made by a succeeding node. In
that case the succeeding node should produce a tuple to return to the
parent Append node from tuples fetched by process_pending_request() when
notified. Repair.
Per buildfarm via Michael Paquier. Back-patch to v14, like the previous
commit.
Thanks to Tom Lane for testing.
Discussion: https://postgr.es/m/YQP0UPT8KmPiHTMs%40paquier.xyz
A pending asynchronous request is handled by process_pending_request(),
which previously not only processed an in-progress remote query but
performed ExecForeignScan() to produce a tuple to return to the local
server asynchronously from the result of the remote query. But that led
to a server crash when executing a query or led to an "InstrStartNode
called twice in a row" or "InstrEndLoop called on running node" failure
when doing EXPLAIN ANALYZE of it, in cases where the plan tree for it
contained multiple async-capable nodes accessing the same
initplan/subplan that contained multiple async-capable nodes scanning
the same foreign tables as for the parent async-capable nodes, as
reported by Andrey Lepikhov. The reason is that the second step in
process_pending_request() invoked when executing the initplan/subplan
for one of the parent async-capable nodes caused recursive execution of
the initplan/subplan for another of the parent async-capable nodes.
To fix, split process_pending_request() into the two steps and postpone
the second step until ForeignAsyncConfigureWait() is called for each of
the pending asynchronous requests. Also, in ExecAppendAsyncEventWait()
we assumed that FDWs would register at least one wait event in a
WaitEventSet created there when they were called from
ForeignAsyncConfigureWait() in that function, but allow FDWs to register
zero wait events in the WaitEventSet; modify ExecAppendAsyncEventWait()
to just return in that case.
Oversight in commit 27e1f1456. Back-patch to v14 where that commit went
in.
Andrey Lepikhov and Etsuro Fujita
Discussion: https://postgr.es/m/fe5eaa19-1704-e4a4-76ee-3b9d37ade399@postgrespro.ru
The point of introducing the hash_mem_multiplier GUC was to let users
reproduce the old behavior of hash aggregation, i.e. that it could use
more than work_mem at need. However, the implementation failed to get
the job done on Win64, where work_mem is clamped to 2GB to protect
various places that calculate memory sizes using "long int". As
written, the same clamp was applied to hash_mem. This resulted in
severe performance regressions for queries requiring a bit more than
2GB for hash aggregation, as they now spill to disk and there's no
way to stop that.
Getting rid of the work_mem restriction seems like a good idea, but
it's a big job and could not conceivably be back-patched. However,
there's only a fairly small number of places that are concerned with
the hash_mem value, and it turns out to be possible to remove the
restriction there without too much code churn or any ABI breaks.
So, let's do that for now to fix the regression, and leave the
larger task for another day.
This patch does introduce a bit more infrastructure that should help
with the larger task, namely pg_bitutils.h support for working with
size_t values.
Per gripe from Laurent Hasson. Back-patch to v13 where the
behavior change came in.
Discussion: https://postgr.es/m/997817.1627074924@sss.pgh.pa.us
Discussion: https://postgr.es/m/MN2PR15MB25601E80A9B6D1BA6F592B1985E39@MN2PR15MB2560.namprd15.prod.outlook.com
Datum sorts can be significantly faster than tuple sorts, especially when
the data type being sorted is a pass-by-value type. Something in the
region of 50-70% performance improvements appear to be possible.
Just in case there's any confusion; the Datum sort is only used when the
targetlist of the Sort node contains a single column, not when there's a
single column in the sort key and multiple items in the target list.
Author: Ronan Dunklau
Reviewed-by: James Coleman, David Rowley, Ranier Vilela, Hou Zhijie
Tested-by: John Naylor
Discussion: https://postgr.es/m/3177670.itZtoPt7T5@aivenronan
Follow-up to 2ed532ee8c, a few error
messages in the logical replication area currently only deal with
tables, but if we're anticipating more relkinds such as sequences
being handled, then these messages also fall into the category
affected by the previous patch, so adjust them too.
Reviewed-by: Michael Paquier <michael@paquier.xyz>
Discussion: https://www.postgresql.org/message-id/c9ba5c6a-4bd5-e12c-1b3c-edbcaedbf392@enterprisedb.com
"Result Cache" was never a great name for this node, but nobody managed
to come up with another name that anyone liked enough. That was until
David Johnston mentioned "Node Memoization", which Tom Lane revised to
just "Memoize". People seem to like "Memoize", so let's do the rename.
Reviewed-by: Justin Pryzby
Discussion: https://postgr.es/m/20210708165145.GG1176@momjian.us
Backpatch-through: 14, where Result Cache was introduced
Similar to 50e17ad28, which allowed hash tables to be used for IN clauses
with a set of constants, here we add the same feature for NOT IN clauses.
NOT IN evaluates the same as: WHERE a <> v1 AND a <> v2 AND a <> v3.
Obviously, if we're using a hash table we must be exactly equivalent to
that and return the same result taking into account that either side of
the condition could contain a NULL. This requires a little bit of
special handling to make work with the hash table version.
When processing NOT IN, the ScalarArrayOpExpr's operator will be the <>
operator. To be able to build and lookup a hash table we must use the
<>'s negator operator. The planner checks if that exists and is hashable
and sets the relevant fields in ScalarArrayOpExpr to instruct the executor
to use hashing.
Author: David Rowley, James Coleman
Reviewed-by: James Coleman, Zhihong Yu
Discussion: https://postgr.es/m/CAApHDvoF1mum_FRk6D621edcB6KSHBi2+GAgWmioj5AhOu2vwQ@mail.gmail.com
Previously, all CustomScan providers had to support projections,
but there may be cases where this is inconvenient. Add a flag
bit to say if it's supported.
Important item for the release notes: this is non-backwards-compatible
since the default is now to assume that CustomScan providers can't
project, instead of assuming that they can. It's fail-soft, but could
result in visible performance penalties due to adding unnecessary
Result nodes.
Sven Klemm, reviewed by Aleksander Alekseev; some cosmetic fiddling
by me.
Discussion: https://postgr.es/m/CAMCrgp1kyakOz6c8aKhNDJXjhQ1dEjEnp+6KNT3KxPrjNtsrDg@mail.gmail.com
Here we alter the code that calls build_pertrans_for_aggref() so that the
function no longer needs to special-case whether it's dealing with an
aggtransfn or an aggcombinefn. This allows us to reuse the
build_aggregate_transfn_expr() function and just get rid of the
build_aggregate_combinefn_expr() completely.
All of the special case code that was in build_pertrans_for_aggref() has
been moved up to the calling functions.
This saves about a dozen lines of code in nodeAgg.c and a few dozen more
in parse_agg.c
Also, rename a few variables in nodeAgg.c to try to make it more clear
that we're working with either a aggtransfn or an aggcombinefn. Some of
the old names would have you believe that we were always working with an
aggtransfn.
Discussion: https://postgr.es/m/CAApHDvptMQ9FmF0D67zC_w88yVnoNVR2+kkOQGUrCmdxWxLULQ@mail.gmail.com
In the "simple Query" code path, it's fine for parse analysis or
execution of a utility statement to scribble on the statement's node
tree, since that'll just be thrown away afterwards. However it's
not fine if the node tree is in the plan cache, as then it'd be
corrupted for subsequent executions. Up to now we've dealt with
that by having individual utility-statement functions apply
copyObject() if they were going to modify the tree. But that's
prone to errors of omission. Bug #17053 from Charles Samborski
shows that CREATE/ALTER DOMAIN didn't get this memo, and can
crash if executed repeatedly from plan cache.
In the back branches, we'll just apply a narrow band-aid for that,
but in HEAD it seems prudent to have a more principled fix that
will close off the possibility of other similar bugs in future.
Hence, let's hoist the responsibility for doing copyObject up into
ProcessUtility from its children, thus ensuring that it happens for
all utility statement types.
Also, modify ProcessUtility's API so that its callers can tell it
whether a copy step is necessary. It turns out that in all cases,
the immediate caller knows whether the node tree is transient, so
this doesn't involve a huge amount of code thrashing. In this way,
while we lose a little bit in the execute-from-cache code path due
to sometimes copying node trees that wouldn't be mutated anyway,
we gain something in the simple-Query code path by not copying
throwaway node trees. Statements that are complex enough to be
expensive to copy are almost certainly ones that would have to be
copied anyway, so the loss in the cache code path shouldn't be much.
(Note that this whole problem applies only to utility statements.
Optimizable statements don't have the issue because we long ago made
the executor treat Plan trees as read-only. Perhaps someday we will
make utility statement execution act likewise, but I'm not holding
my breath.)
Discussion: https://postgr.es/m/931771.1623893989@sss.pgh.pa.us
Discussion: https://postgr.es/m/17053-3ca3f501bbc212b4@postgresql.org
Commit b676ac443b optimized handling of tuple slots with bulk inserts
into foreign tables, so that the slots are initialized only once and
reused for all batches. The data was however copied into the slots only
after the initialization, inserting duplicate values when the slot gets
reused. Fixed by moving the ExecCopySlot outside the init branch.
The existing postgres_fdw tests failed to catch this due to inserting
data into foreign tables without unique indexes, and then checking only
the number of inserted rows. This adds a new test with both a unique
index and a check of inserted values.
Reported-by: Alexander Pyhalov
Discussion: https://postgr.es/m/7a8cf8d56b3d18e5c0bccd6cd42d04ac%40postgrespro.ru
Commit b663a41363 introduced bulk inserts for FDW, but the handling of
tuple slots turned out to be problematic for two reasons. Firstly, the
slots were re-created for each individual batch. Secondly, all slots
referenced the same tuple descriptor - with reasonably small batches
this is not an issue, but with large batches this triggers O(N^2)
behavior in the resource owner code.
These two issues work against each other - to reduce the number of times
a slot has to be created/dropped, larger batches are needed. However,
the larger the batch, the more expensive the resource owner gets. For
practical batch sizes (100 - 1000) this would not be a big problem, as
the benefits (latency savings) greatly exceed the resource owner costs.
But for extremely large batches it might be much worse, possibly even
losing with non-batching mode.
Fixed by initializing tuple slots only once (and reusing them across
batches) and by using a new tuple descriptor copy for each slot.
Discussion: https://postgr.es/m/ebbbcc7d-4286-8c28-0272-61b4753af761%40enterprisedb.com
We've accumulated quite a mix of instances of "an SQL" and "a SQL" in the
documents. It would be good to be a bit more consistent with these.
The most recent version of the SQL standard I looked at seems to prefer
"an SQL". That seems like a good lead to follow, so here we change all
instances of "a SQL" to become "an SQL". Most instances correctly use
"an SQL" already, so it also makes sense to use the dominant variation in
order to minimise churn.
Additionally, there were some other abbreviations that needed to be
adjusted. FSM, SSPI, SRF and a few others. Also fix some pronounceable,
abbreviations to use "a" instead of "an". For example, "a SASL" instead
of "an SASL".
Here I've only adjusted the documents and error messages. Many others
still exist in source code comments. Translator hint comments seem to be
the biggest culprit. It currently does not seem worth the churn to change
these.
Discussion: https://postgr.es/m/CAApHDvpML27UqFXnrYO1MJddsKVMQoiZisPvsAGhKE_tsKXquw%40mail.gmail.com
Commit 2453ea142 redefined pg_proc.proargtypes to include the types of
OUT parameters, for procedures only. While that had some advantages
for implementing the SQL-spec behavior of DROP PROCEDURE, it was pretty
disastrous from a number of other perspectives. Notably, since the
primary key of pg_proc is name + proargtypes, this made it possible to
have multiple procedures with identical names + input arguments and
differing output argument types. That would make it impossible to call
any one of the procedures by writing just NULL (or "?", or any other
data-type-free notation) for the output argument(s). The change also
seems likely to cause grave confusion for client applications that
examine pg_proc and expect the traditional definition of proargtypes.
Hence, revert the definition of proargtypes to what it was, and
undo a number of complications that had been added to support that.
To support the SQL-spec behavior of DROP PROCEDURE, when there are
no argmode markers in the command's parameter list, we perform the
lookup both ways (that is, matching against both proargtypes and
proallargtypes), succeeding if we get just one unique match.
In principle this could result in ambiguous-function failures
that would not happen when using only one of the two rules.
However, overloading of procedure names is thought to be a pretty
rare usage, so this shouldn't cause many problems in practice.
Postgres-specific code such as pg_dump can defend against any
possibility of such failures by being careful to specify argmodes
for all procedure arguments.
This also fixes a few other bugs in the area of CALL statements
with named parameters, and improves the documentation a little.
catversion bump forced because the representation of procedures
with OUT arguments changes.
Discussion: https://postgr.es/m/3742981.1621533210@sss.pgh.pa.us
Previously, we left the EPQ sub-executor alone until ExecEndLockRows.
This caused any buffer pins or other resources that it might hold to
remain held until ExecutorEnd, which in some code paths means that
they are held till the Portal is closed. That can cause user-visible
problems, such as blocking VACUUM; and it's unlike the behavior of
ordinary table-scanning nodes, which will have released all buffer
pins by the time they return an EOF indication.
We can make LockRows work more like other plan nodes by calling
EvalPlanQualEnd just before returning NULL. We still need to call it
in ExecEndLockRows in case the node was not run to completion, but in
the normal case the second call does nothing and costs little.
Per report from Yura Sokolov. In principle this is a longstanding
bug, but in view of the lack of other complaints and the low severity
of the consequences, I chose not to back-patch.
Discussion: https://postgr.es/m/4aa370cb91ecf2f9885d98b80ad1109c@postgrespro.ru
In cases where run-time pruning isn't required, the synchronous and
asynchronous subplans for an async-aware Append node determined using
classify_matching_subplans() should be re-used when rescanning the node,
but the previous code re-determined them using that function repeatedly
each time when rescanning the node, leading to incorrect results in a
normal build and an Assert failure in an Assert-enabled build as that
function doesn't assume that it's called repeatedly in such cases. Fix
the code as mentioned above.
My oversight in commit 27e1f1456.
While at it, initialize async-related pointers/variables to NULL/zero
explicitly in ExecInitAppend() and ExecReScanAppend(), just to be sure.
(The variables would have been set to zero before we get to the latter
function, but let's do so.)
Reviewed-by: Kyotaro Horiguchi
Discussion: https://postgr.es/m/CAPmGK16Q4B2_KY%2BJH7rb7wQbw54AUprp7TMekGTd2T1B62yysQ%40mail.gmail.com
When the planner considered using a Result Cache node to cache results
from the inner side of a Nested Loop Join, it failed to consider that the
inner path's parameterization may not be the entire join condition. If
the join was marked as inner_unique then we may accidentally put the cache
in singlerow mode. This meant that entries would be marked as complete
after caching the first row. That was wrong as if only part of the join
condition was parameterized then the uniqueness of the unique join was not
guaranteed at the Result Cache's level. The uniqueness is only guaranteed
after Nested Loop applies the join filter. If subsequent rows were found,
this would lead to:
ERROR: cache entry already complete
This could have been fixed by only putting the cache in singlerow mode if
the entire join condition was parameterized. However, Nested Loop will
only read its inner side so far as the first matching row when the join is
unique, so that might mean we never get an opportunity to mark cache
entries as complete. Since non-complete cache entries are useless for
subsequent lookups, we just don't bother considering a Result Cache path
in this case.
In passing, remove the XXX comment that claimed the above ERROR might be
better suited to be an Assert. After there being an actual case which
triggered it, it seems better to keep it an ERROR.
Reported-by: David Christensen
Discussion: https://postgr.es/m/CAOxo6X+dy-V58iEPFgst8ahPKEU+38NZzUuc+a7wDBZd4TrHMQ@mail.gmail.com
We consider this supported (though I've got my doubts that it's a
good idea, because tableoid is not immutable). However, several
code paths failed to fill the field in soon enough, causing such
a GENERATED expression to see zero or the wrong value. This
occurred when ALTER TABLE adds a new GENERATED column to a table
with existing rows, and during regular INSERT or UPDATE on a
foreign table with GENERATED columns.
Noted during investigation of a report from Vitaly Ustinov.
Back-patch to v12 where GENERATED came in.
Discussion: https://postgr.es/m/CAM_DEiWR2DPT6U4xb-Ehigozzd3n3G37ZB1+867zbsEVtYoJww@mail.gmail.com
COMMIT/ROLLBACK necessarily destroys all snapshots within the session.
The original implementation of intra-procedure transactions just
cavalierly did that, ignoring the fact that this left us executing in
a rather different environment than normal. In particular, it turns
out that handling of toasted datums depends rather critically on there
being an outer ActiveSnapshot: otherwise, when SPI or the core
executor pop whatever snapshot they used and return, it's unsafe to
dereference any toasted datums that may appear in the query result.
It's possible to demonstrate "no known snapshots" and "missing chunk
number N for toast value" errors as a result of this oversight.
Historically this outer snapshot has been held by the Portal code,
and that seems like a good plan to preserve. So add infrastructure
to pquery.c to allow re-establishing the Portal-owned snapshot if it's
not there anymore, and add enough bookkeeping support that we can tell
whether it is or not.
We can't, however, just re-establish the Portal snapshot as part of
COMMIT/ROLLBACK. As in normal transaction start, acquiring the first
snapshot should wait until after SET and LOCK commands. Hence, teach
spi.c about doing this at the right time. (Note that this patch
doesn't fix the problem for any PLs that try to run intra-procedure
transactions without using SPI to execute SQL commands.)
This makes SPI's no_snapshots parameter rather a misnomer, so in HEAD,
rename that to allow_nonatomic.
replication/logical/worker.c also needs some fixes, because it wasn't
careful to hold a snapshot open around AFTER trigger execution.
That code doesn't use a Portal, which I suspect someday we're gonna
have to fix. But for now, just rearrange the order of operations.
This includes back-patching the recent addition of finish_estate()
to centralize the cleanup logic there.
This also back-patches commit 2ecfeda3e into v13, to improve the
test coverage for worker.c (it was that test that exposed that
worker.c's snapshot management is wrong).
Per bug #15990 from Andreas Wicht. Back-patch to v11 where
intra-procedure COMMIT was added.
Discussion: https://postgr.es/m/15990-eee2ac466b11293d@postgresql.org
This seems to be leftover from ea15e1867 and from when we used to evaluate
SRFs at each node.
Since there is an unconditional "return" at the end of the loop body, only
1 loop is ever possible, so we can just change this into an if condition.
There is no actual bug being fixed here so no back-patch. It seems fine to
just fix this anomaly in master only.
Author: Greg Nancarrow
Discussion: https://postgr.es/m/CAJcOf-d7T1q0az-D8evWXnsuBZjigT04WkV5hCAOEJQZRWy28w@mail.gmail.com
Also "make reformat-dat-files".
The only change worthy of note is that pgindent messed up the formatting
of launcher.c's struct LogicalRepWorkerId, which led me to notice that
that struct wasn't used at all anymore, so I just took it out.
EXPLAIN ANALYZE for an async-capable ForeignScan node associated with
postgres_fdw is done just by using instrumentation for ExecProcNode()
called from the node's callbacks, causing the following problems:
1) If the remote table to scan is empty, the node is incorrectly
considered as "never executed" by the command even if the node is
executed, as ExecProcNode() isn't called from the node's callbacks at
all in that case.
2) The command fails to collect timings for things other than
ExecProcNode() done in the node, such as creating a cursor for the
node's remote query.
To fix these problems, add instrumentation for async-capable nodes, and
modify postgres_fdw accordingly.
My oversight in commit 27e1f1456.
While at it, update a comment for the AsyncRequest struct in execnodes.h
and the documentation for the ForeignAsyncRequest API in fdwhandler.sgml
to match the code in ExecAsyncAppendResponse() in nodeAppend.c, and fix
typos in comments in nodeAppend.c.
Per report from Andrey Lepikhov, though I didn't use his patch.
Reviewed-by: Andrey Lepikhov
Discussion: https://postgr.es/m/2eb662bb-105d-fc20-7412-2f027cc3ca72%40postgrespro.ru
It's unusual to have any resjunk columns in an ON CONFLICT ... UPDATE
list, but it can happen when MULTIEXPR_SUBLINK SubPlans are present.
If it happens, the ON CONFLICT UPDATE code path would end up storing
tuples that include the values of the extra resjunk columns. That's
fairly harmless in the short run, but if new columns are added to
the table then the values would become accessible, possibly leading
to malfunctions if they don't match the datatypes of the new columns.
This had escaped notice through a confluence of missing sanity checks,
including
* There's no cross-check that a tuple presented to heap_insert or
heap_update matches the table rowtype. While it's difficult to
check that fully at reasonable cost, we can easily add assertions
that there aren't too many columns.
* The output-column-assignment cases in execExprInterp.c lacked
any sanity checks on the output column numbers, which seems like
an oversight considering there are plenty of assertion checks on
input column numbers. Add assertions there too.
* We failed to apply nodeModifyTable's ExecCheckPlanOutput() to
the ON CONFLICT UPDATE tlist. That wouldn't have caught this
specific error, since that function is chartered to ignore resjunk
columns; but it sure seems like a bad omission now that we've seen
this bug.
In HEAD, the right way to fix this is to make the processing of
ON CONFLICT UPDATE tlists work the same as regular UPDATE tlists
now do, that is don't add "SET x = x" entries, and use
ExecBuildUpdateProjection to evaluate the tlist and combine it with
old values of the not-set columns. This adds a little complication
to ExecBuildUpdateProjection, but allows removal of a comparable
amount of now-dead code from the planner.
In the back branches, the most expedient solution seems to be to
(a) use an output slot for the ON CONFLICT UPDATE projection that
actually matches the target table, and then (b) invent a variant of
ExecBuildProjectionInfo that can be told to not store values resulting
from resjunk columns, so it doesn't try to store into nonexistent
columns of the output slot. (We can't simply ignore the resjunk columns
altogether; they have to be evaluated for MULTIEXPR_SUBLINK to work.)
This works back to v10. In 9.6, projections work much differently and
we can't cheaply give them such an option. The 9.6 version of this
patch works by inserting a JunkFilter when it's necessary to get rid
of resjunk columns.
In addition, v11 and up have the reverse problem when trying to
perform ON CONFLICT UPDATE on a partitioned table. Through a
further oversight, adjust_partition_tlist() discarded resjunk columns
when re-ordering the ON CONFLICT UPDATE tlist to match a partition.
This accidentally prevented the storing-bogus-tuples problem, but
at the cost that MULTIEXPR_SUBLINK cases didn't work, typically
crashing if more than one row has to be updated. Fix by preserving
resjunk columns in that routine. (I failed to resist the temptation
to add more assertions there too, and to do some minor code
beautification.)
Per report from Andres Freund. Back-patch to all supported branches.
Security: CVE-2021-32028
While we were (mostly) careful about ensuring that the dimensions of
arrays aren't large enough to cause integer overflow, the lower bound
values were generally not checked. This allows situations where
lower_bound + dimension overflows an integer. It seems that that's
harmless so far as array reading is concerned, except that array
elements with subscripts notionally exceeding INT_MAX are inaccessible.
However, it confuses various array-assignment logic, resulting in a
potential for memory stomps.
Fix by adding checks that array lower bounds aren't large enough to
cause lower_bound + dimension to overflow. (Note: this results in
disallowing cases where the last subscript position would be exactly
INT_MAX. In principle we could probably allow that, but there's a lot
of code that computes lower_bound + dimension and would need adjustment.
It seems doubtful that it's worth the trouble/risk to allow it.)
Somewhat independently of that, array_set_element() was careless
about possible overflow when checking the subscript of a fixed-length
array, creating a different route to memory stomps. Fix that too.
Security: CVE-2021-32027
In 9eacee2e6, I included some code to verify the cache's memory tracking
is correct by counting up the number of entries and the memory they use
each time we evict something from the cache. Those values are then
compared to the expected values using Assert. The problem is that this
requires looping over the entire cache hash table each time we evict an
entry from the cache. That can be pretty expensive, as noted by Pavel
Stehule.
Here we move this memory accounting checking code so that we only verify
it on cassert builds once when shutting down the Result Cache node.
Aside from the performance increase, this has two distinct advantages:
1) We do the memory checks at the last possible moment before destroying
the cache. This means we'll now catch accounting problems that might
sneak in after a cache eviction.
2) We now do the memory Assert checks when there were no cache evictions.
This increases the coverage.
One small disadvantage is that we'll now miss any memory tracking issues
that somehow managed to resolve themselves by the end of execution.
However, it seems to me that such a memory tracking problem would be quite
unlikely, and likely somewhat less harmful if one were to exist.
In passing, adjust the loop over the hash table to use the standard
simplehash.h method of iteration.
Reported-by: Pavel Stehule
Discussion: https://postgr.es/m/CAFj8pRAzgoSkdEiqrKbT=7yG9FA5fjUAP3jmJywuDqYq6Ki5ug@mail.gmail.com
This is cleanup for commit 27e1f1456:
* ExecAppendAsyncEventWait(), which was modified a bit further by commit
a8af856d3, duplicated the same nevents calculation. Simplify the code
a little bit to avoid the duplication. Update comments there.
* Add an assertion to ExecAppendAsyncRequest().
* Update a comment about merging the async_capable options from input
relations in merge_fdw_options(), per complaint from Kyotaro Horiguchi.
* Add a comment for fetch_more_data_begin().
Author: Etsuro Fujita
Discussion: https://postgr.es/m/CAPmGK1637W30Wx3MnrReewhafn6F_0J76mrJGoFXFnpPq4QfvA%40mail.gmail.com
Adopt a more consistent policy about what slot-type-specific
getsysattr functions should do when system attributes are not
available. To wit, they should all throw the same user-oriented
error, rather than variously crashing or emitting developer-oriented
messages.
This closes a identifiable problem in commits a71cfc56b and
3fb93103a (in v13 and v12), so back-patch into those branches,
along with a test case to try to ensure we don't break it again.
It is not known that any of the former crash cases are reachable
in HEAD, but this seems like a good safety improvement in any case.
Discussion: https://postgr.es/m/141051591267657@mail.yandex.ru
During queries coming from ri_triggers.c, we need to omit partitions
that are marked pending detach -- otherwise, the RI query is tricked
into allowing a row into the referencing table whose corresponding row
is in the detached partition. Which is bogus: once the detach operation
completes, the row becomes an orphan.
However, the code was not doing that in repeatable-read transactions,
because relcache kept a copy of the partition descriptor that included
the partition, and used it in the RI query. This commit changes the
partdesc cache code to only keep descriptors that aren't dependent on
a snapshot (namely: those where no detached partition exist, and those
where detached partitions are included). When a partdesc-without-
detached-partitions is requested, we create one afresh each time; also,
those partdescs are stored in PortalContext instead of
CacheMemoryContext.
find_inheritance_children gets a new output *detached_exist boolean,
which indicates whether any partition marked pending-detach is found.
Its "include_detached" input flag is changed to "omit_detached", because
that name captures desired the semantics more naturally.
CreatePartitionDirectory() and RelationGetPartitionDesc() arguments are
identically renamed.
This was noticed because a buildfarm member that runs with relcache
clobbering, which would not keep the improperly cached partdesc, broke
one test, which led us to realize that the expected output of that test
was bogus. This commit also corrects that expected output.
Author: Amit Langote <amitlangote09@gmail.com>
Author: Álvaro Herrera <alvherre@alvh.no-ip.org>
Discussion: https://postgr.es/m/3269784.1617215412@sss.pgh.pa.us
Previously, it was pg_stat_activity.queryid to match the
pg_stat_statements queryid column. This is an adjustment to patch
4f0b0966c8. This also adjusts some of the internal function calls to
match. Catversion bumped.
Reported-by: Álvaro Herrera, Julien Rouhaud
Discussion: https://postgr.es/m/20210408032704.GA7498@alvherre.pgsql
Commit e717a9a18 changed the longstanding rule that prosrc is NOT NULL
because when a SQL-language function is written in SQL-standard style,
we don't currently have anything useful to put there. This seems a poor
decision though, as it could easily have negative impacts on external
PLs (opening them to crashes they didn't use to have, for instance).
SQL-function-related code can just as easily test "is prosqlbody not
null" as "is prosrc null", so there's no real gain there either.
Hence, revert the NOT NULL marking removal and adjust related logic.
For now, we just put an empty string into prosrc for SQL-standard
functions. Maybe we'll have a better idea later, although the
history of things like pg_attrdef.adsrc suggests that it's not
easy to maintain a string equivalent of a node tree.
This also adds an assertion that queryDesc->sourceText != NULL
to standard_ExecutorStart. We'd been silently relying on that
for awhile, so let's make it less silent.
Also fix some overlooked documentation and test cases.
Discussion: https://postgr.es/m/2197698.1617984583@sss.pgh.pa.us
Previously, get_cached_rowtype() cached a pointer to a reference-counted
tuple descriptor from the typcache, relying on the ExprContextCallback
mechanism to release the tupdesc refcount when the expression tree
using the tupdesc was destroyed. This worked fine when it was designed,
but the introduction of within-DO-block COMMITs broke it. The refcount
is logged in a transaction-lifespan resource owner, but plpgsql won't
destroy simple expressions made within the DO block (before its first
commit) until the DO block is exited. That results in a warning about
a leaked tupdesc refcount when the COMMIT destroys the original resource
owner, and then an error about the active resource owner not holding a
matching refcount when the expression is destroyed.
To fix, get rid of the need to have a shutdown callback at all, by
instead caching a pointer to the relevant typcache entry. Those
survive for the life of the backend, so we needn't worry about the
pointer becoming stale. (For registered RECORD types, we can still
cache a pointer to the tupdesc, knowing that it won't change for the
life of the backend.) This mechanism has been in use in plpgsql
and expandedrecord.c since commit 4b93f5799, and seems to work well.
This change requires modifying the ExprEvalStep structs used by the
relevant expression step types, which is slightly worrisome for
back-patching. However, there seems no good reason for extensions
to be familiar with the details of these particular sub-structs.
Per report from Rohit Bhogate. Back-patch to v11 where within-DO-block
COMMITs became a thing.
Discussion: https://postgr.es/m/CAAV6ZkQRCVBh8qAY+SZiHnz+U+FqAGBBDaDTjF2yiKa2nJSLKg@mail.gmail.com
Ignore parallel workers in pg_stat_statements
Oversight in 4f0b0966c8 which exposed queryid in parallel workers.
Counters are aggregated by the main backend process so parallel workers
would report duplicated activity, and could also report activity for the
wrong entry as they are only aware of the top level queryid.
Fix thinko in pg_stat_get_activity when retrieving the queryid.
Remove unnecessary call to pgstat_report_queryid().
Reported-by: Amit Kapila, Andres Freund, Thomas Munro
Discussion: https://postgr.es/m/20210408051735.lfbdzun5zdlax5gd@alap3.anarazel.dep634GTSOqnDW86Owrn6qDAVosC5dJjXjp7BMfc5Gz1Q@mail.gmail.com
Author: Julien Rouhaud
ScalarArrayOpExprs with "useOr=true" and a set of Consts on the righthand
side have traditionally been evaluated by using a linear search over the
array. When these arrays contain large numbers of elements then this
linear search could become a significant part of execution time.
Here we add a new method of evaluating ScalarArrayOpExpr expressions to
allow them to be evaluated by first building a hash table containing each
element, then on subsequent evaluations, we just probe that hash table to
determine if there is a match.
The planner is in charge of determining when this optimization is possible
and it enables it by setting hashfuncid in the ScalarArrayOpExpr. The
executor will only perform the hash table evaluation when the hashfuncid
is set.
This means that not all cases are optimized. For example CHECK constraints
containing an IN clause won't go through the planner, so won't get the
hashfuncid set. We could maybe do something about that at some later
date. The reason we're not doing it now is from fear that we may slow
down cases where the expression is evaluated only once. Those cases can
be common, for example, a single row INSERT to a table with a CHECK
constraint containing an IN clause.
In the planner, we enable this when there are suitable hash functions for
the ScalarArrayOpExpr's operator and only when there is at least
MIN_ARRAY_SIZE_FOR_HASHED_SAOP elements in the array. The threshold is
currently set to 9.
Author: James Coleman, David Rowley
Reviewed-by: David Rowley, Tomas Vondra, Heikki Linnakangas
Discussion: https://postgr.es/m/CAAaqYe8x62+=wn0zvNKCj55tPpg-JBHzhZFFc6ANovdqFw7-dA@mail.gmail.com
It's far from clear that this is the right approach - but a good
portion of the buildfarm has been red for a few hours, on the last day
of the CF. And this fixes at least the obvious crash. So let's go with
that for now.
Discussion: https://postgr.es/m/20210407225806.majgznh4lk34hjvu%40alap3.anarazel.de
This adds support for writing CREATE FUNCTION and CREATE PROCEDURE
statements for language SQL with a function body that conforms to the
SQL standard and is portable to other implementations.
Instead of the PostgreSQL-specific AS $$ string literal $$ syntax,
this allows writing out the SQL statements making up the body
unquoted, either as a single statement:
CREATE FUNCTION add(a integer, b integer) RETURNS integer
LANGUAGE SQL
RETURN a + b;
or as a block
CREATE PROCEDURE insert_data(a integer, b integer)
LANGUAGE SQL
BEGIN ATOMIC
INSERT INTO tbl VALUES (a);
INSERT INTO tbl VALUES (b);
END;
The function body is parsed at function definition time and stored as
expression nodes in a new pg_proc column prosqlbody. So at run time,
no further parsing is required.
However, this form does not support polymorphic arguments, because
there is no more parse analysis done at call time.
Dependencies between the function and the objects it uses are fully
tracked.
A new RETURN statement is introduced. This can only be used inside
function bodies. Internally, it is treated much like a SELECT
statement.
psql needs some new intelligence to keep track of function body
boundaries so that it doesn't send off statements when it sees
semicolons that are inside a function body.
Tested-by: Jaime Casanova <jcasanov@systemguards.com.ec>
Reviewed-by: Julien Rouhaud <rjuju123@gmail.com>
Discussion: https://www.postgresql.org/message-id/flat/1c11f1eb-f00c-43b7-799d-2d44132c02d7@2ndquadrant.com
Use the in-core query id computation for pg_stat_activity,
log_line_prefix, and EXPLAIN VERBOSE.
Similar to other fields in pg_stat_activity, only the queryid from the
top level statements are exposed, and if the backends status isn't
active then the queryid from the last executed statements is displayed.
Add a %Q placeholder to include the queryid in log_line_prefix, which
will also only expose top level statements.
For EXPLAIN VERBOSE, if a query identifier has been computed, either by
enabling compute_query_id or using a third-party module, display it.
Bump catalog version.
Discussion: https://postgr.es/m/20210407125726.tkvjdbw76hxnpwfi@nol
Author: Julien Rouhaud
Reviewed-by: Alvaro Herrera, Nitin Jadhav, Zhihong Yu
Delay creation of the projections for INSERT and UPDATE tuples
until they're needed. This saves a pretty fair amount of work
when only some of the partitions are actually touched.
The logic associated with identifying junk columns in UPDATE/DELETE
is moved to another loop, allowing removal of one loop over the
target relations; but it didn't actually change at all.
Extracted from a larger patch, which seemed to me to be too messy
to push in one commit.
Amit Langote, reviewed at different times by Heikki Linnakangas and
myself
Discussion: https://postgr.es/m/CA+HiwqG7ZruBmmih3wPsBZ4s0H2EhywrnXEduckY5Hr3fWzPWA@mail.gmail.com
Arrange to do some things on-demand, rather than immediately during
executor startup, because there's a fair chance of never having to do
them at all:
* Don't open result relations' indexes until needed.
* Don't initialize partition tuple routing, nor the child-to-root
tuple conversion map, until needed.
This wins in UPDATEs on partitioned tables when only some of the
partitions will actually receive updates; with larger partition
counts the savings is quite noticeable. Also, we can remove some
sketchy heuristics in ExecInitModifyTable about whether to set up
tuple routing.
Also, remove execPartition.c's private hash table tracking which
partitions were already opened by the ModifyTable node. Instead
use the hash added to ModifyTable itself by commit 86dc90056.
To allow lazy computation of the conversion maps, we now set
ri_RootResultRelInfo in all child ResultRelInfos. We formerly set it
only in some, not terribly well-defined, cases. This has user-visible
side effects in that now more error messages refer to the root
relation instead of some partition (and provide error data in the
root's column order, too). It looks to me like this is a strict
improvement in consistency, so I don't have a problem with the
output changes visible in this commit.
Extracted from a larger patch, which seemed to me to be too messy
to push in one commit.
Amit Langote, reviewed at different times by Heikki Linnakangas and
myself
Discussion: https://postgr.es/m/CA+HiwqG7ZruBmmih3wPsBZ4s0H2EhywrnXEduckY5Hr3fWzPWA@mail.gmail.com
Andrew Gierth reported that it's possible to crash the backend if no
pg_attrdef record is found to match an attribute that has atthasdef set.
AttrDefaultFetch warns about this situation, but then leaves behind
a relation tupdesc that has null "adbin" pointer(s), which most places
don't guard against.
We considered promoting the warning to an error, but throwing errors
during relcache load is pretty drastic: it effectively locks one out
of using the relation at all. What seems better is to leave the
load-time behavior as a warning, but then throw an error in any code
path that wants to use a default and can't find it. This confines
the error to a subset of INSERT/UPDATE operations on the table, and
in particular will at least allow a pg_dump to succeed.
Also, we should fix AttrDefaultFetch to not leave any null pointers
in the tupdesc, because that just creates an untested bug hazard.
While at it, apply the same philosophy of "warn at load, throw error
only upon use of the known-missing info" to CHECK constraints.
CheckConstraintFetch is very nearly the same logic as AttrDefaultFetch,
but for reasons lost in the mists of time, it was throwing ERROR for
the same cases that AttrDefaultFetch treats as WARNING. Make the two
functions more nearly alike.
In passing, get rid of potentially-O(N^2) loops in equalTupleDesc
by making AttrDefaultFetch sort the entries after fetching them,
so that equalTupleDesc can assume that entries in two equal tupdescs
must be in matching order. (CheckConstraintFetch already was sorting
CHECK constraints, but equalTupleDesc hadn't been told about it.)
There's some argument for back-patching this, but with such a small
number of field reports, I'm content to fix it in HEAD.
Discussion: https://postgr.es/m/87pmzaq4gx.fsf@news-spur.riddles.org.uk
Adjust the number of events given to WaitEventSetWait() so that it
doesn't exceed the maximum number of events in the WaitEventSet given
to that function (set->nevents_space) in hopes of making the buildfarm
green.
Per valgrind failure report from Tom Lane and the buildfarm.
Author: Etsuro Fujita
Discussion: https://postgr.es/m/3411577.1617289776%40sss.pgh.pa.us
Testing if an unsigned variable is >= 0 is pretty pointless.
There's likely enough code in remove_cache_entry() to verify the cache
memory accounting is correct in assert enabled builds. These Asserts
were not adding much extra cover, even if they had been checking >= 0 on a
signed variable.
Reported-by: Andres Freund
Discussion: https://postgr.es/m/20210402204734.6mo3nfacnljlicgn@alap3.anarazel.de
Here we add a new executor node type named "Result Cache". The planner
can include this node type in the plan to have the executor cache the
results from the inner side of parameterized nested loop joins. This
allows caching of tuples for sets of parameters so that in the event that
the node sees the same parameter values again, it can just return the
cached tuples instead of rescanning the inner side of the join all over
again. Internally, result cache uses a hash table in order to quickly
find tuples that have been previously cached.
For certain data sets, this can significantly improve the performance of
joins. The best cases for using this new node type are for join problems
where a large portion of the tuples from the inner side of the join have
no join partner on the outer side of the join. In such cases, hash join
would have to hash values that are never looked up, thus bloating the hash
table and possibly causing it to multi-batch. Merge joins would have to
skip over all of the unmatched rows. If we use a nested loop join with a
result cache, then we only cache tuples that have at least one join
partner on the outer side of the join. The benefits of using a
parameterized nested loop with a result cache increase when there are
fewer distinct values being looked up and the number of lookups of each
value is large. Also, hash probes to lookup the cache can be much faster
than the hash probe in a hash join as it's common that the result cache's
hash table is much smaller than the hash join's due to result cache only
caching useful tuples rather than all tuples from the inner side of the
join. This variation in hash probe performance is more significant when
the hash join's hash table no longer fits into the CPU's L3 cache, but the
result cache's hash table does. The apparent "random" access of hash
buckets with each hash probe can cause a poor L3 cache hit ratio for large
hash tables. Smaller hash tables generally perform better.
The hash table used for the cache limits itself to not exceeding work_mem
* hash_mem_multiplier in size. We maintain a dlist of keys for this cache
and when we're adding new tuples and realize we've exceeded the memory
budget, we evict cache entries starting with the least recently used ones
until we have enough memory to add the new tuples to the cache.
For parameterized nested loop joins, we now consider using one of these
result cache nodes in between the nested loop node and its inner node. We
determine when this might be useful based on cost, which is primarily
driven off of what the expected cache hit ratio will be. Estimating the
cache hit ratio relies on having good distinct estimates on the nested
loop's parameters.
For now, the planner will only consider using a result cache for
parameterized nested loop joins. This works for both normal joins and
also for LATERAL type joins to subqueries. It is possible to use this new
node for other uses in the future. For example, to cache results from
correlated subqueries. However, that's not done here due to some
difficulties obtaining a distinct estimation on the outer plan to
calculate the estimated cache hit ratio. Currently we plan the inner plan
before planning the outer plan so there is no good way to know if a result
cache would be useful or not since we can't estimate the number of times
the subplan will be called until the outer plan is generated.
The functionality being added here is newly introducing a dependency on
the return value of estimate_num_groups() during the join search.
Previously, during the join search, we only ever needed to perform
selectivity estimations. With this commit, we need to use
estimate_num_groups() in order to estimate what the hit ratio on the
result cache will be. In simple terms, if we expect 10 distinct values
and we expect 1000 outer rows, then we'll estimate the hit ratio to be
99%. Since cache hits are very cheap compared to scanning the underlying
nodes on the inner side of the nested loop join, then this will
significantly reduce the planner's cost for the join. However, it's
fairly easy to see here that things will go bad when estimate_num_groups()
incorrectly returns a value that's significantly lower than the actual
number of distinct values. If this happens then that may cause us to make
use of a nested loop join with a result cache instead of some other join
type, such as a merge or hash join. Our distinct estimations have been
known to be a source of trouble in the past, so the extra reliance on them
here could cause the planner to choose slower plans than it did previous
to having this feature. Distinct estimations are also fairly hard to
estimate accurately when several tables have been joined already or when a
WHERE clause filters out a set of values that are correlated to the
expressions we're estimating the number of distinct value for.
For now, the costing we perform during query planning for result caches
does put quite a bit of faith in the distinct estimations being accurate.
When these are accurate then we should generally see faster execution
times for plans containing a result cache. However, in the real world, we
may find that we need to either change the costings to put less trust in
the distinct estimations being accurate or perhaps even disable this
feature by default. There's always an element of risk when we teach the
query planner to do new tricks that it decides to use that new trick at
the wrong time and causes a regression. Users may opt to get the old
behavior by turning the feature off using the enable_resultcache GUC.
Currently, this is enabled by default. It remains to be seen if we'll
maintain that setting for the release.
Additionally, the name "Result Cache" is the best name I could think of
for this new node at the time I started writing the patch. Nobody seems
to strongly dislike the name. A few people did suggest other names but no
other name seemed to dominate in the brief discussion that there was about
names. Let's allow the beta period to see if the current name pleases
enough people. If there's some consensus on a better name, then we can
change it before the release. Please see the 2nd discussion link below
for the discussion on the "Result Cache" name.
Author: David Rowley
Reviewed-by: Andy Fan, Justin Pryzby, Zhihong Yu, Hou Zhijie
Tested-By: Konstantin Knizhnik
Discussion: https://postgr.es/m/CAApHDvrPcQyQdWERGYWx8J%2B2DLUNgXu%2BfOSbQ1UscxrunyXyrQ%40mail.gmail.com
Discussion: https://postgr.es/m/CAApHDvq=yQXr5kqhRviT2RhNKwToaWr9JAN5t+5_PzhuRJ3wvg@mail.gmail.com
This removes "Add Result Cache executor node". It seems that something
weird is going on with the tracking of cache hits and misses as
highlighted by many buildfarm animals. It's not yet clear what the
problem is as other parts of the plan indicate that the cache did work
correctly, it's just the hits and misses that were being reported as 0.
This is especially a bad time to have the buildfarm so broken, so
reverting before too many more animals go red.
Discussion: https://postgr.es/m/CAApHDvq_hydhfovm4=izgWs+C5HqEeRScjMbOgbpC-jRAeK3Yw@mail.gmail.com
Here we add a new executor node type named "Result Cache". The planner
can include this node type in the plan to have the executor cache the
results from the inner side of parameterized nested loop joins. This
allows caching of tuples for sets of parameters so that in the event that
the node sees the same parameter values again, it can just return the
cached tuples instead of rescanning the inner side of the join all over
again. Internally, result cache uses a hash table in order to quickly
find tuples that have been previously cached.
For certain data sets, this can significantly improve the performance of
joins. The best cases for using this new node type are for join problems
where a large portion of the tuples from the inner side of the join have
no join partner on the outer side of the join. In such cases, hash join
would have to hash values that are never looked up, thus bloating the hash
table and possibly causing it to multi-batch. Merge joins would have to
skip over all of the unmatched rows. If we use a nested loop join with a
result cache, then we only cache tuples that have at least one join
partner on the outer side of the join. The benefits of using a
parameterized nested loop with a result cache increase when there are
fewer distinct values being looked up and the number of lookups of each
value is large. Also, hash probes to lookup the cache can be much faster
than the hash probe in a hash join as it's common that the result cache's
hash table is much smaller than the hash join's due to result cache only
caching useful tuples rather than all tuples from the inner side of the
join. This variation in hash probe performance is more significant when
the hash join's hash table no longer fits into the CPU's L3 cache, but the
result cache's hash table does. The apparent "random" access of hash
buckets with each hash probe can cause a poor L3 cache hit ratio for large
hash tables. Smaller hash tables generally perform better.
The hash table used for the cache limits itself to not exceeding work_mem
* hash_mem_multiplier in size. We maintain a dlist of keys for this cache
and when we're adding new tuples and realize we've exceeded the memory
budget, we evict cache entries starting with the least recently used ones
until we have enough memory to add the new tuples to the cache.
For parameterized nested loop joins, we now consider using one of these
result cache nodes in between the nested loop node and its inner node. We
determine when this might be useful based on cost, which is primarily
driven off of what the expected cache hit ratio will be. Estimating the
cache hit ratio relies on having good distinct estimates on the nested
loop's parameters.
For now, the planner will only consider using a result cache for
parameterized nested loop joins. This works for both normal joins and
also for LATERAL type joins to subqueries. It is possible to use this new
node for other uses in the future. For example, to cache results from
correlated subqueries. However, that's not done here due to some
difficulties obtaining a distinct estimation on the outer plan to
calculate the estimated cache hit ratio. Currently we plan the inner plan
before planning the outer plan so there is no good way to know if a result
cache would be useful or not since we can't estimate the number of times
the subplan will be called until the outer plan is generated.
The functionality being added here is newly introducing a dependency on
the return value of estimate_num_groups() during the join search.
Previously, during the join search, we only ever needed to perform
selectivity estimations. With this commit, we need to use
estimate_num_groups() in order to estimate what the hit ratio on the
result cache will be. In simple terms, if we expect 10 distinct values
and we expect 1000 outer rows, then we'll estimate the hit ratio to be
99%. Since cache hits are very cheap compared to scanning the underlying
nodes on the inner side of the nested loop join, then this will
significantly reduce the planner's cost for the join. However, it's
fairly easy to see here that things will go bad when estimate_num_groups()
incorrectly returns a value that's significantly lower than the actual
number of distinct values. If this happens then that may cause us to make
use of a nested loop join with a result cache instead of some other join
type, such as a merge or hash join. Our distinct estimations have been
known to be a source of trouble in the past, so the extra reliance on them
here could cause the planner to choose slower plans than it did previous
to having this feature. Distinct estimations are also fairly hard to
estimate accurately when several tables have been joined already or when a
WHERE clause filters out a set of values that are correlated to the
expressions we're estimating the number of distinct value for.
For now, the costing we perform during query planning for result caches
does put quite a bit of faith in the distinct estimations being accurate.
When these are accurate then we should generally see faster execution
times for plans containing a result cache. However, in the real world, we
may find that we need to either change the costings to put less trust in
the distinct estimations being accurate or perhaps even disable this
feature by default. There's always an element of risk when we teach the
query planner to do new tricks that it decides to use that new trick at
the wrong time and causes a regression. Users may opt to get the old
behavior by turning the feature off using the enable_resultcache GUC.
Currently, this is enabled by default. It remains to be seen if we'll
maintain that setting for the release.
Additionally, the name "Result Cache" is the best name I could think of
for this new node at the time I started writing the patch. Nobody seems
to strongly dislike the name. A few people did suggest other names but no
other name seemed to dominate in the brief discussion that there was about
names. Let's allow the beta period to see if the current name pleases
enough people. If there's some consensus on a better name, then we can
change it before the release. Please see the 2nd discussion link below
for the discussion on the "Result Cache" name.
Author: David Rowley
Reviewed-by: Andy Fan, Justin Pryzby, Zhihong Yu
Tested-By: Konstantin Knizhnik
Discussion: https://postgr.es/m/CAApHDvrPcQyQdWERGYWx8J%2B2DLUNgXu%2BfOSbQ1UscxrunyXyrQ%40mail.gmail.com
Discussion: https://postgr.es/m/CAApHDvq=yQXr5kqhRviT2RhNKwToaWr9JAN5t+5_PzhuRJ3wvg@mail.gmail.com
This patch makes two closely related sets of changes:
1. For UPDATE, the subplan of the ModifyTable node now only delivers
the new values of the changed columns (i.e., the expressions computed
in the query's SET clause) plus row identity information such as CTID.
ModifyTable must re-fetch the original tuple to merge in the old
values of any unchanged columns. The core advantage of this is that
the changed columns are uniform across all tables of an inherited or
partitioned target relation, whereas the other columns might not be.
A secondary advantage, when the UPDATE involves joins, is that less
data needs to pass through the plan tree. The disadvantage of course
is an extra fetch of each tuple to be updated. However, that seems to
be very nearly free in context; even worst-case tests don't show it to
add more than a couple percent to the total query cost. At some point
it might be interesting to combine the re-fetch with the tuple access
that ModifyTable must do anyway to mark the old tuple dead; but that
would require a good deal of refactoring and it seems it wouldn't buy
all that much, so this patch doesn't attempt it.
2. For inherited UPDATE/DELETE, instead of generating a separate
subplan for each target relation, we now generate a single subplan
that is just exactly like a SELECT's plan, then stick ModifyTable
on top of that. To let ModifyTable know which target relation a
given incoming row refers to, a tableoid junk column is added to
the row identity information. This gets rid of the horrid hack
that was inheritance_planner(), eliminating O(N^2) planning cost
and memory consumption in cases where there were many unprunable
target relations.
Point 2 of course requires point 1, so that there is a uniform
definition of the non-junk columns to be returned by the subplan.
We can't insist on uniform definition of the row identity junk
columns however, if we want to keep the ability to have both
plain and foreign tables in a partitioning hierarchy. Since
it wouldn't scale very far to have every child table have its
own row identity column, this patch includes provisions to merge
similar row identity columns into one column of the subplan result.
In particular, we can merge the whole-row Vars typically used as
row identity by FDWs into one column by pretending they are type
RECORD. (It's still okay for the actual composite Datums to be
labeled with the table's rowtype OID, though.)
There is more that can be done to file down residual inefficiencies
in this patch, but it seems to be committable now.
FDW authors should note several API changes:
* The argument list for AddForeignUpdateTargets() has changed, and so
has the method it must use for adding junk columns to the query. Call
add_row_identity_var() instead of manipulating the parse tree directly.
You might want to reconsider exactly what you're adding, too.
* PlanDirectModify() must now work a little harder to find the
ForeignScan plan node; if the foreign table is part of a partitioning
hierarchy then the ForeignScan might not be the direct child of
ModifyTable. See postgres_fdw for sample code.
* To check whether a relation is a target relation, it's no
longer sufficient to compare its relid to root->parse->resultRelation.
Instead, check it against all_result_relids or leaf_result_relids,
as appropriate.
Amit Langote and Tom Lane
Discussion: https://postgr.es/m/CA+HiwqHpHdqdDn48yCEhynnniahH78rwcrv1rEX65-fsZGBOLQ@mail.gmail.com
This implements asynchronous execution, which runs multiple parts of a
non-parallel-aware Append concurrently rather than serially to improve
performance when possible. Currently, the only node type that can be
run concurrently is a ForeignScan that is an immediate child of such an
Append. In the case where such ForeignScans access data on different
remote servers, this would run those ForeignScans concurrently, and
overlap the remote operations to be performed simultaneously, so it'll
improve the performance especially when the operations involve
time-consuming ones such as remote join and remote aggregation.
We may extend this to other node types such as joins or aggregates over
ForeignScans in the future.
This also adds the support for postgres_fdw, which is enabled by the
table-level/server-level option "async_capable". The default is false.
Robert Haas, Kyotaro Horiguchi, Thomas Munro, and myself. This commit
is mostly based on the patch proposed by Robert Haas, but also uses
stuff from the patch proposed by Kyotaro Horiguchi and from the patch
proposed by Thomas Munro. Reviewed by Kyotaro Horiguchi, Konstantin
Knizhnik, Andrey Lepikhov, Movead Li, Thomas Munro, Justin Pryzby, and
others.
Discussion: https://postgr.es/m/CA%2BTgmoaXQEt4tZ03FtQhnzeDEMzBck%2BLrni0UWHVVgOTnA6C1w%40mail.gmail.com
Discussion: https://postgr.es/m/CA%2BhUKGLBRyu0rHrDCMC4%3DRn3252gogyp1SjOgG8SEKKZv%3DFwfQ%40mail.gmail.com
Discussion: https://postgr.es/m/20200228.170650.667613673625155850.horikyota.ntt%40gmail.com
Allow a partition be detached from its partitioned table without
blocking concurrent queries, by running in two transactions and only
requiring ShareUpdateExclusive in the partitioned table.
Because it runs in two transactions, it cannot be used in a transaction
block. This is the main reason to use dedicated syntax: so that users
can choose to use the original mode if they need it. But also, it
doesn't work when a default partition exists (because an exclusive lock
would still need to be obtained on it, in order to change its partition
constraint.)
In case the second transaction is cancelled or a crash occurs, there's
ALTER TABLE .. DETACH PARTITION .. FINALIZE, which executes the final
steps.
The main trick to make this work is the addition of column
pg_inherits.inhdetachpending, initially false; can only be set true in
the first part of this command. Once that is committed, concurrent
transactions that use a PartitionDirectory will include or ignore
partitions so marked: in optimizer they are ignored if the row is marked
committed for the snapshot; in executor they are always included. As a
result, and because of the way PartitionDirectory caches partition
descriptors, queries that were planned before the detach will see the
rows in the detached partition and queries that are planned after the
detach, won't.
A CHECK constraint is created that duplicates the partition constraint.
This is probably not strictly necessary, and some users will prefer to
remove it afterwards, but if the partition is re-attached to a
partitioned table, the constraint needn't be rechecked.
Author: Álvaro Herrera <alvherre@alvh.no-ip.org>
Reviewed-by: Amit Langote <amitlangote09@gmail.com>
Reviewed-by: Justin Pryzby <pryzby@telsasoft.com>
Discussion: https://postgr.es/m/20200803234854.GA24158@alvherre.pgsql
To allow inserts in parallel-mode this feature has to ensure that all the
constraints, triggers, etc. are parallel-safe for the partition hierarchy
which is costly and we need to find a better way to do that. Additionally,
we could have used existing cached information in some cases like indexes,
domains, etc. to determine the parallel-safety.
List of commits reverted, in reverse chronological order:
ed62d3737c Doc: Update description for parallel insert reloption.
c8f78b6161 Add a new GUC and a reloption to enable inserts in parallel-mode.
c5be48f092 Improve FK trigger parallel-safety check added by 05c8482f7f.
e2cda3c20a Fix use of relcache TriggerDesc field introduced by commit 05c8482f7f.
e4e87a32cc Fix valgrind issue in commit 05c8482f7f.
05c8482f7f Enable parallel SELECT for "INSERT INTO ... SELECT ...".
Discussion: https://postgr.es/m/E1lMiB9-0001c3-SY@gemulon.postgresql.org
Commit 3048898e dropped -ING from some wait event names that correspond
to barrier phases. Update the phases' names to match.
While we're here making cosmetic changes, also rename "DONE" to "FREE".
That pairs better with "ALLOCATE", and describes the activity that
actually happens in that phase (as we do for the other phases) rather
than describing a state. The distinction is clearer after bugfix commit
3b8981b6 split the phase into two. As for the growth barriers, rename
their "ALLOCATE" phase to "REALLOCATE", which is probably a better
description of what happens then. Also improve the comments about
the phases a bit.
Discussion: https://postgr.es/m/CA%2BhUKG%2BMDpwF2Eo2LAvzd%3DpOh81wUTsrwU1uAwR-v6OGBB6%2B7g%40mail.gmail.com
With very unlucky timing and parallel_leader_participation off, PHJ
could attempt to access per-batch state just as it was being freed.
There was code intended to prevent that by checking for a cleared
pointer, but it was buggy.
Fix, by introducing an extra barrier phase. The new phase
PHJ_BUILD_RUNNING means that it's safe to access the per-batch state to
find a batch to help with, and PHJ_BUILD_DONE means that it is too late.
The last to detach will free the array of per-batch state as before, but
now it will also atomically advance the phase at the same time, so that
late attachers can avoid the hazard, without the data race. This
mirrors the way per-batch hash tables are freed (see phases
PHJ_BATCH_PROBING and PHJ_BATCH_DONE).
Revealed by a one-off build farm failure, where BarrierAttach() failed a
sanity check assertion, because the memory had been clobbered by
dsa_free().
Back-patch to 11, where the code arrived.
Reported-by: Michael Paquier <michael@paquier.xyz>
Discussion: https://postgr.es/m/20200929061142.GA29096%40paquier.xyz
Parallel SELECT can't be utilized for INSERT in the following cases:
- INSERT statement uses the ON CONFLICT DO UPDATE clause
- Target table has a parallel-unsafe: trigger, index expression or
predicate, column default expression or check constraint
- Target table has a parallel-unsafe domain constraint on any column
- Target table is a partitioned table with a parallel-unsafe partition key
expression or support function
The planner is updated to perform additional parallel-safety checks for
the cases listed above, for determining whether it is safe to run INSERT
in parallel-mode with an underlying parallel SELECT. The planner will
consider using parallel SELECT for "INSERT INTO ... SELECT ...", provided
nothing unsafe is found from the additional parallel-safety checks, or
from the existing parallel-safety checks for SELECT.
While checking parallel-safety, we need to check it for all the partitions
on the table which can be costly especially when we decide not to use a
parallel plan. So, in a separate patch, we will introduce a GUC and or a
reloption to enable/disable parallelism for Insert statements.
Prior to entering parallel-mode for the execution of INSERT with parallel
SELECT, a TransactionId is acquired and assigned to the current
transaction state. This is necessary to prevent the INSERT from attempting
to assign the TransactionId whilst in parallel-mode, which is not allowed.
This approach has a disadvantage in that if the underlying SELECT does not
return any rows, then the TransactionId is not used, however that
shouldn't happen in practice in many cases.
Author: Greg Nancarrow, Amit Langote, Amit Kapila
Reviewed-by: Amit Langote, Hou Zhijie, Takayuki Tsunakawa, Antonin Houska, Bharath Rupireddy, Dilip Kumar, Vignesh C, Zhihong Yu, Amit Kapila
Tested-by: Tang, Haiying
Discussion: https://postgr.es/m/CAJcOf-cXnB5cnMKqWEp2E2z7Mvcd04iLVmV=qpFJrR3AcrTS3g@mail.gmail.com
Discussion: https://postgr.es/m/CAJcOf-fAdj=nDKMsRhQzndm-O13NY4dL6xGcEvdX5Xvbbi0V7g@mail.gmail.com
Throw a "function protocol violation" error if a function in FROM
tries to return a set though it wasn't marked proretset. Although
such cases work at the moment, it doesn't seem like something we
want to guarantee will keep working. Besides, there are other
negative consequences of not setting the proretset flag, such as
potentially bad plans.
No back-patch, since if there is any third-party code violating
this expectation, people wouldn't appreciate us breaking it in
a minor release.
Discussion: https://postgr.es/m/1636062.1615141782@sss.pgh.pa.us
This adds a new executor node named TID Range Scan. The query planner
will generate paths for TID Range scans when quals are discovered on base
relations which search for ranges on the table's ctid column. These
ranges may be open at either end. For example, WHERE ctid >= '(10,0)';
will return all tuples on page 10 and over.
To support this, two new optional callback functions have been added to
table AM. scan_set_tidrange is used to set the scan range to just the
given range of TIDs. scan_getnextslot_tidrange fetches the next tuple
in the given range.
For AMs were scanning ranges of TIDs would not make sense, these functions
can be set to NULL in the TableAmRoutine. The query planner won't
generate TID Range Scan Paths in that case.
Author: Edmund Horner, David Rowley
Reviewed-by: David Rowley, Tomas Vondra, Tom Lane, Andres Freund, Zhihong Yu
Discussion: https://postgr.es/m/CAMyN-kB-nFTkF=VA_JPwFNo08S0d-Yk0F741S2B7LDmYAi8eyA@mail.gmail.com
A cross-partition update on a partitioned table is implemented as a
delete followed by an insert. With foreign partitions, this was however
causing issues, because the FDW and core may disagree on when to enable
batching. postgres_fdw was only allowing batching for plain inserts
(CMD_INSERT) while core was trying to batch the insert component of the
cross-partition update. Fix by restricting core to apply batching only
to plain CMD_INSERT queries.
It's possible to allow batching for cross-partition updates, but that
will require more extensive changes, so better to leave that for a
separate patch.
Author: Amit Langote
Reviewed-by: Tomas Vondra, Takayuki Tsunakawa
Discussion: https://postgr.es/m/20200628151002.7x5laxwpgvkyiu3q@development
The inner loop in switchToPresortedPrefixMode() can be implemented
as a conventional integer-counter for() loop, removing a couple of
redundant boolean state variables. The old logic here was a remnant
of earlier development, but as things now stand there's no reason
for extra complexity.
Also, annotate the test case added by 82e0e2930 to explain why it
manages to hit the corner case fixed in that commit, and add an
EXPLAIN to verify that it's creating an incremental-sort plan.
Back-patch to v13, like the previous patch.
James Coleman and Tom Lane
Discussion: https://postgr.es/m/16846-ae49f51ac379a4cb@postgresql.org
If ExecGetInsertedCols(), ExecGetUpdatedCols() or ExecGetExtraUpdatedCols()
were called with a ResultRelInfo that's not in the range table and isn't a
partition routing target, the functions would dereference a NULL pointer,
relinfo->ri_RootResultRelInfo. Such ResultRelInfos are created when firing
RI triggers in tables that are not modified directly. None of the current
callers of these functions pass such relations, so this isn't a live bug,
but let's make them more robust.
Also update comment in ResultRelInfo; after commit 6214e2b228,
ri_RangeTableIndex is zero for ResultRelInfos created for partition tuple
routing.
Noted by Coverity. Backpatch down to v11, like commit 6214e2b228.
Reviewed-by: Tom Lane, Amit Langote
If a cross-partition UPDATE violates a constraint on the target partition,
and the columns in the new partition are in different physical order than
in the parent, the error message can reveal columns that the user does not
have SELECT permission on. A similar bug was fixed earlier in commit
804b6b6db4.
The cause of the bug is that the callers of the
ExecBuildSlotValueDescription() function got confused when constructing
the list of modified columns. If the tuple was routed from a parent, we
converted the tuple to the parent's format, but the list of modified
columns was grabbed directly from the child's RTE entry.
ExecUpdateLockMode() had a similar issue. That lead to confusion on which
columns are key columns, leading to wrong tuple lock being taken on tables
referenced by foreign keys, when a row is updated with INSERT ON CONFLICT
UPDATE. A new isolation test is added for that corner case.
With this patch, the ri_RangeTableIndex field is no longer set for
partitions that don't have an entry in the range table. Previously, it was
set to the RTE entry of the parent relation, but that was confusing.
NOTE: This modifies the ResultRelInfo struct, replacing the
ri_PartitionRoot field with ri_RootResultRelInfo. That's a bit risky to
backpatch, because it breaks any extensions accessing the field. The
change that ri_RangeTableIndex is not set for partitions could potentially
break extensions, too. The ResultRelInfos are visible to FDWs at least,
and this patch required small changes to postgres_fdw. Nevertheless, this
seem like the least bad option. I don't think these fields widely used in
extensions; I don't think there are FDWs out there that uses the FDW
"direct update" API, other than postgres_fdw. If there is, you will get a
compilation error, so hopefully it is caught quickly.
Backpatch to 11, where support for both cross-partition UPDATEs, and unique
indexes on partitioned tables, were added.
Reviewed-by: Amit Langote
Security: CVE-2021-3393
Commit 230230223 taught nodeAgg.c that, when spilling tuples from
memory in an oversized hash aggregation, it only needed to spill
input columns referenced in the node's tlist and quals. Unfortunately,
that's wrong: we also have to save the grouping columns. The error
is masked in common cases because the grouping columns also appear
in the tlist, but that's not necessarily true. The main category
of plans where it's not true seem to come from semijoins ("WHERE
outercol IN (SELECT innercol FROM innertable)") where the innercol
needs an implicit promotion to make it comparable to the outercol.
The grouping column will be "innercol::promotedtype", but that
expression appears nowhere in the Agg node's own tlist and quals;
only the bare "innercol" is found in the tlist.
I spent quite a bit of time looking for a suitable regression test
case for this, without much success. If the number of distinct
values of the innercol is large enough to make spilling happen,
the planner tends to prefer a non-HashAgg plan, at least for
problem sizes that are reasonable to use in the regression tests.
So, no new regression test. However, this patch does demonstrably
fix the originally-reported test case.
Per report from s.p.e (at) gmx-topmail.de. Backpatch to v13
where the troublesome code came in.
Discussion: https://postgr.es/m/trinity-1c565d44-159f-488b-a518-caf13883134f-1611835701633@3c-app-gmx-bap78
switchToPresortedPrefixMode() did the wrong thing if it detected
a batch boundary just at the last tuple of a fullsort group.
The initially-reported symptom was a "retrieved too many tuples in a
bounded sort" error, but the test case added here just silently gives
the wrong answer without this patch.
I (tgl) am not really happy about committing this patch without review
from the incremental-sort authors, but they seem AWOL and we are hard
against a release deadline. This does demonstrably make some cases
better, anyway.
Per bug #16846 from Yoran Heling. Back-patch to v13 where incremental
sort was introduced.
Neil Chen
Discussion: https://postgr.es/m/16846-ae49f51ac379a4cb@postgresql.org
perform_pruning_combine_step() was not taught about the number of
partition indexes used in hash partitioning; more embarrassingly,
get_matching_hash_bounds() also had it wrong. These errors are masked
in the common case where all the partitions have the same modulus
and no partition is missing. However, with missing or unequal-size
partitions, we could erroneously prune some partitions that need
to be scanned, leading to silently wrong query answers.
While a minimal-footprint fix for this could be to export
get_partition_bound_num_indexes and make the incorrect functions use it,
I'm of the opinion that that function should never have existed in the
first place. It's not reasonable data structure design that
PartitionBoundInfoData lacks any explicit record of the length of
its indexes[] array. Perhaps that was all right when it could always
be assumed equal to ndatums, but something should have been done about
it as soon as that stopped being true. Putting in an explicit
"nindexes" field makes both partition_bounds_equal() and
partition_bounds_copy() simpler, safer, and faster than before,
and removes explicit knowledge of the number-of-partition-indexes
rules from some other places too.
This change also makes get_hash_partition_greatest_modulus obsolete.
I left that in place in case any external code uses it, but no core
code does anymore.
Per bug #16840 from Michał Albrycht. Back-patch to v11 where the
hash partitioning code came in. (In the back branches, add the new
field at the end of PartitionBoundInfoData to minimize ABI risks.)
Discussion: https://postgr.es/m/16840-571a22976f829ad4@postgresql.org
When building aggregate expression steps, strict checks need a bailout
jump for when a null value is encountered, so there is a list of steps
that require later adjustment. Adding entries to that list for steps
that aren't actually strict would be harmless, except that there is an
Assert which catches them. This leads to spurious errors on asserts
builds, for data sets that trigger parallel aggregation of an
aggregate with a non-strict deserialization function (no such
aggregates exist in the core system).
Repair by not adding the adjustment entry when it's not needed.
Backpatch back to 11 where the code was introduced.
Per a report from Darafei (Komzpa) of the PostGIS project; analysis
and patch by me.
Discussion: https://postgr.es/m/87mty7peb3.fsf@news-spur.riddles.org.uk
SPI_execute_with_receiver and SPI_cursor_parse_open_with_paramlist are
new in v14 (cf. commit 2f48ede08). Before they can get out the door,
let's change their APIs to follow the practice recently established by
SPI_prepare_extended etc: shove all optional arguments into a struct
that callers are supposed to pre-zero. The hope is to allow future
addition of more options without either API breakage or a continuing
proliferation of new SPI entry points. With that in mind, choose
slightly more generic names for them: SPI_execute_extended and
SPI_cursor_parse_open respectively.
Discussion: https://postgr.es/m/CAFj8pRCLPdDAETvR7Po7gC5y_ibkn_-bOzbeJb39WHms01194Q@mail.gmail.com
This patch essentially is cleaning up technical debt left behind
by the original implementation of plpgsql procedures, particularly
commit d92bc83c4. That patch (or more precisely, follow-on patches
fixing its worst bugs) forced us to re-plan CALL and DO statements
each time through, if we're in a non-atomic context. That wasn't
for any fundamental reason, but just because use of a saved plan
requires having a ResourceOwner to hold a reference count for the
plan, and we had no suitable resowner at hand, nor would the
available APIs support using one if we did. While it's not that
expensive to create a "plan" for CALL/DO, the cycles do add up
in repeated executions.
This patch therefore makes the following API changes:
* GetCachedPlan/ReleaseCachedPlan are modified to let the caller
specify which resowner to use to pin the plan, rather than forcing
use of CurrentResourceOwner.
* spi.c gains a "SPI_execute_plan_extended" entry point that lets
callers say which resowner to use to pin the plan. This borrows the
idea of an options struct from the recently added SPI_prepare_extended,
hopefully allowing future options to be added without more API breaks.
This supersedes SPI_execute_plan_with_paramlist (which I've marked
deprecated) as well as SPI_execute_plan_with_receiver (which is new
in v14, so I just took it out altogether).
* I also took the opportunity to remove the crude hack of letting
plpgsql reach into SPI private data structures to mark SPI plans as
"no_snapshot". It's better to treat that as an option of
SPI_prepare_extended.
Now, when running a non-atomic procedure or DO block that contains
any CALL or DO commands, plpgsql creates a ResourceOwner that
will be used to pin the plans of the CALL/DO commands. (In an
atomic context, we just use CurrentResourceOwner, as before.)
Having done this, we can just save CALL/DO plans normally,
whether or not they are used across transaction boundaries.
This seems to be good for something like 2X speedup of a CALL
of a trivial procedure with a few simple argument expressions.
By restricting the creation of an extra ResourceOwner like this,
there's essentially zero penalty in cases that can't benefit.
Pavel Stehule, with some further hacking by me
Discussion: https://postgr.es/m/CAFj8pRCLPdDAETvR7Po7gC5y_ibkn_-bOzbeJb39WHms01194Q@mail.gmail.com
ExecInitModifyTable has to initialize batching for all result relations,
not just the first one. Furthermore, when junk filters were necessary,
the pointer pointed past the mtstate->resultRelInfo array.
Per reports from multiple non-x86 animals (florican, locust, ...).
Discussion: https://postgr.es/m/20200628151002.7x5laxwpgvkyiu3q@development
Extends the FDW API to allow batching inserts into foreign tables. That
is usually much more efficient than inserting individual rows, due to
high latency for each round-trip to the foreign server.
It was possible to implement something similar in the regular FDW API,
but it was inconvenient and there were issues with reporting the number
of actually inserted rows etc. This extends the FDW API with two new
functions:
* GetForeignModifyBatchSize - allows the FDW picking optimal batch size
* ExecForeignBatchInsert - inserts a batch of rows at once
Currently, only INSERT queries support batching. Support for DELETE and
UPDATE may be added in the future.
This also implements batching for postgres_fdw. The batch size may be
specified using "batch_size" option both at the server and table level.
The initial patch version was written by me, but it was rewritten and
improved in many ways by Takayuki Tsunakawa.
Author: Takayuki Tsunakawa
Reviewed-by: Tomas Vondra, Amit Langote
Discussion: https://postgr.es/m/20200628151002.7x5laxwpgvkyiu3q@development
Somebody extended search_plan_tree() to treat MergeAppend exactly
like Append, which is 100% wrong, because unlike Append we can't
assume that only one input node is actively returning tuples.
Hence a cursor using a MergeAppend across a UNION ALL or inheritance
tree could falsely match a WHERE CURRENT OF query at a row that
isn't actually the cursor's current output row, but coincidentally
has the same TID (in a different table) as the current output row.
Delete the faulty code; this means that such a case will now return
an error like 'cursor "foo" is not a simply updatable scan of table
"bar"', instead of silently misbehaving. Users should not find that
surprising though, as the same cursor query could have failed that way
already depending on the chosen plan. (It would fail like that if the
sort were done with an explicit Sort node instead of MergeAppend.)
Expand the clearly-inadequate commentary to be more explicit about
what this code is doing, in hopes of forestalling future mistakes.
It's been like this for awhile, so back-patch to all supported
branches.
Discussion: https://postgr.es/m/482865.1611075182@sss.pgh.pa.us
execCurrent.c's search_plan_tree() assumed that ForeignScanStates
and CustomScanStates necessarily have a valid ss_currentRelation.
This is demonstrably untrue for postgres_fdw's remote join and
remote aggregation plans, and non-leaf custom scans might not have
an identifiable scan relation either. Avoid crashing by ignoring
such nodes when the field is null.
This solution will lead to errors like 'cursor "foo" is not a
simply updatable scan of table "bar"' in cases where maybe we
could have allowed WHERE CURRENT OF to work. That's not an issue
for postgres_fdw's usages, since joins or aggregations would render
WHERE CURRENT OF invalid anyway. But an otherwise-transparent
upper level custom scan node might find this annoying. When and if
someone cares to expend work on such a scenario, we could invent a
custom-scan-provider callback to determine what's safe.
Report and patch by David Geier, commentary by me. It's been like
this for awhile, so back-patch to all supported branches.
Discussion: https://postgr.es/m/0253344d-9bdd-11c4-7f0d-d88c02cd7991@swarm64.com
Add an executor aminsert() hint mechanism that informs index AMs that
the incoming index tuple (the tuple that accompanies the hint) is not
being inserted by execution of an SQL statement that logically modifies
any of the index's key columns.
The hint is received by indexes when an UPDATE takes place that does not
apply an optimization like heapam's HOT (though only for indexes where
all key columns are logically unchanged). Any index tuple that receives
the hint on insert is expected to be a duplicate of at least one
existing older version that is needed for the same logical row. Related
versions will typically be stored on the same index page, at least
within index AMs that apply the hint.
Recognizing the difference between MVCC version churn duplicates and
true logical row duplicates at the index AM level can help with cleanup
of garbage index tuples. Cleanup can intelligently target tuples that
are likely to be garbage, without wasting too many cycles on less
promising tuples/pages (index pages with little or no version churn).
This is infrastructure for an upcoming commit that will teach nbtree to
perform bottom-up index deletion. No index AM actually applies the hint
just yet.
Author: Peter Geoghegan <pg@bowt.ie>
Reviewed-By: Victor Yegorov <vyegorov@gmail.com>
Discussion: https://postgr.es/m/CAH2-Wz=CEKFa74EScx_hFVshCOn6AA5T-ajFASTdzipdkLTNQQ@mail.gmail.com
Invent new RawParseModes that allow the core grammar to handle
pl/pgsql expressions and assignments directly, and thereby get rid
of a lot of hackery in pl/pgsql's parser. This moves a good deal
of knowledge about pl/pgsql into the core code: notably, we have to
invent a CoercionContext that matches pl/pgsql's (rather dubious)
historical behavior for assignment coercions. That's getting away
from the original idea of pl/pgsql as an arm's-length extension of
the core, but really we crossed that bridge a long time ago.
The main advantage of doing this is that we can now use the core
parser to generate FieldStore and/or SubscriptingRef nodes to handle
assignments to pl/pgsql variables that are records or arrays. That
fixes a number of cases that had never been implemented in pl/pgsql
assignment, such as nested records and array slicing, and it allows
pl/pgsql assignment to support the datatype-specific subscripting
behaviors introduced in commit c7aba7c14.
There are cosmetic benefits too: when a syntax error occurs in a
pl/pgsql expression, the error report no longer includes the confusing
"SELECT" keyword that used to get prefixed to the expression text.
Also, there seem to be some small speed gains.
Discussion: https://postgr.es/m/4165684.1607707277@sss.pgh.pa.us
This patch essentially allows gram.y to implement a family of related
syntax trees, rather than necessarily always parsing a list of SQL
statements. raw_parser() gains a new argument, enum RawParseMode,
to say what to do. As proof of concept, add a mode that just parses
a TypeName without any other decoration, and use that to greatly
simplify typeStringToTypeName().
In addition, invent a new SPI entry point SPI_prepare_extended() to
allow SPI users (particularly plpgsql) to get at this new functionality.
In hopes of making this the last variant of SPI_prepare(), set up its
additional arguments as a struct rather than direct arguments, and
promise that future additions to the struct can default to zero.
SPI_prepare_cursor() and SPI_prepare_params() can perhaps go away at
some point.
Discussion: https://postgr.es/m/4165684.1607707277@sss.pgh.pa.us
Before processing tuples, agg_refill_hash_table() was setting all
pergroup pointers to NULL to signal to advance_aggregates() that it
should not attempt to advance groups that had spilled.
The problem was that it also set the pergroups for sorted grouping
sets to NULL, which caused rescanning to fail.
Instead, change agg_refill_hash_table() to only set the pergroups for
hashed grouping sets to NULL; and when compiling the expression, pass
doSort=false.
Reported-by: Alexander Lakhin
Discussion: https://postgr.es/m/16784-7ff169bf2c3d1588%40postgresql.org
Backpatch-through: 13
Multiranges are basically sorted arrays of non-overlapping ranges with
set-theoretic operations defined over them.
Since v14, each range type automatically gets a corresponding multirange
datatype. There are both manual and automatic mechanisms for naming multirange
types. Once can specify multirange type name using multirange_type_name
attribute in CREATE TYPE. Otherwise, a multirange type name is generated
automatically. If the range type name contains "range" then we change that to
"multirange". Otherwise, we add "_multirange" to the end.
Implementation of multiranges comes with a space-efficient internal
representation format, which evades extra paddings and duplicated storage of
oids. Altogether this format allows fetching a particular range by its index
in O(n).
Statistic gathering and selectivity estimation are implemented for multiranges.
For this purpose, stored multirange is approximated as union range without gaps.
This field will likely need improvements in the future.
Catversion is bumped.
Discussion: https://postgr.es/m/CALNJ-vSUpQ_Y%3DjXvTxt1VYFztaBSsWVXeF1y6gTYQ4bOiWDLgQ%40mail.gmail.com
Discussion: https://postgr.es/m/a0b8026459d1e6167933be2104a6174e7d40d0ab.camel%40j-davis.com#fe7218c83b08068bfffb0c5293eceda0
Author: Paul Jungwirth, revised by me
Reviewed-by: David Fetter, Corey Huinker, Jeff Davis, Pavel Stehule
Reviewed-by: Alvaro Herrera, Tom Lane, Isaac Morland, David G. Johnston
Reviewed-by: Zhihong Yu, Alexander Korotkov
Invent a new flag bit HASH_STRINGS to specify C-string hashing, which
was formerly the default; and add assertions insisting that exactly
one of the bits HASH_STRINGS, HASH_BLOBS, and HASH_FUNCTION be set.
This is in hopes of preventing recurrences of the type of oversight
fixed in commit a1b8aa1e4 (i.e., mistakenly omitting HASH_BLOBS).
Also, when HASH_STRINGS is specified, insist that the keysize be
more than 8 bytes. This is a heuristic, but it should catch
accidental use of HASH_STRINGS for integer or pointer keys.
(Nearly all existing use-cases set the keysize to NAMEDATALEN or
more, so there's little reason to think this restriction should
be problematic.)
Tweak hash_create() to insist that the HASH_ELEM flag be set, and
remove the defaults it had for keysize and entrysize. Since those
defaults were undocumented and basically useless, no callers
omitted HASH_ELEM anyway.
Also, remove memset's zeroing the HASHCTL parameter struct from
those callers that had one. This has never been really necessary,
and while it wasn't a bad coding convention it was confusing that
some callers did it and some did not. We might as well save a few
cycles by standardizing on "not".
Also improve the documentation for hash_create().
In passing, improve reinit.c's usage of a hash table by storing
the key as a binary Oid rather than a string; and, since that's
a temporary hash table, allocate it in CurrentMemoryContext for
neatness.
Discussion: https://postgr.es/m/590625.1607878171@sss.pgh.pa.us
Commit c7aba7c14 didn't add this, but after more fooling with the
feature I feel that it'd be useful. To make this possible, refactor
getSubscriptingRoutines() so that the caller is responsible for
throwing any error. (In clauses.c, I just chose to make the
most conservative assumption rather than throwing an error. We don't
expect failures there anyway really, so the code space for an error
message would be a poor investment.)
This patch generalizes the subscripting infrastructure so that any
data type can be subscripted, if it provides a handler function to
define what that means. Traditional variable-length (varlena) arrays
all use array_subscript_handler(), while the existing fixed-length
types that support subscripting use raw_array_subscript_handler().
It's expected that other types that want to use subscripting notation
will define their own handlers. (This patch provides no such new
features, though; it only lays the foundation for them.)
To do this, move the parser's semantic processing of subscripts
(including coercion to whatever data type is required) into a
method callback supplied by the handler. On the execution side,
replace the ExecEvalSubscriptingRef* layer of functions with direct
calls to callback-supplied execution routines. (Thus, essentially
no new run-time overhead should be caused by this patch. Indeed,
there is room to remove some overhead by supplying specialized
execution routines. This patch does a little bit in that line,
but more could be done.)
Additional work is required here and there to remove formerly
hard-wired assumptions about the result type, collation, etc
of a SubscriptingRef expression node; and to remove assumptions
that the subscript values must be integers.
One useful side-effect of this is that we now have a less squishy
mechanism for identifying whether a data type is a "true" array:
instead of wiring in weird rules about typlen, we can look to see
if pg_type.typsubscript == F_ARRAY_SUBSCRIPT_HANDLER. For this
to be bulletproof, we have to forbid user-defined types from using
that handler directly; but there seems no good reason for them to
do so.
This patch also removes assumptions that the number of subscripts
is limited to MAXDIM (6), or indeed has any hard-wired limit.
That limit still applies to types handled by array_subscript_handler
or raw_array_subscript_handler, but to discourage other dependencies
on this constant, I've moved it from c.h to utils/array.h.
Dmitry Dolgov, reviewed at various times by Tom Lane, Arthur Zakirov,
Peter Eisentraut, Pavel Stehule
Discussion: https://postgr.es/m/CA+q6zcVDuGBv=M0FqBYX8DPebS3F_0KQ6OVFobGJPM507_SZ_w@mail.gmail.com
Discussion: https://postgr.es/m/CA+q6zcVovR+XY4mfk-7oNk-rF91gH0PebnNfuUjuuDsyHjOcVA@mail.gmail.com
currtid() and currtid2() are an undocumented set of functions whose sole
known user is the Postgres ODBC driver, able to retrieve the latest TID
version for a tuple given by the caller of those functions.
As used by Postgres ODBC, currtid() is a shortcut able to retrieve the
last TID loaded into a backend by passing an OID of 0 (magic value)
after a tuple insertion. This is removed in this commit, as it became
obsolete after the driver began using "RETURNING ctid" with inserts, a
clause supported since Postgres 8.2 (using RETURNING is better for
performance anyway as it reduces the number of round-trips to the
backend).
currtid2() is still used by the driver, so this remains around for now.
Note that this function is kept in its original shape for backward
compatibility reasons.
Per discussion with many people, including Andres Freund, Peter
Eisentraut, Álvaro Herrera, Hiroshi Inoue, Tom Lane and myself.
Bump catalog version.
Discussion: https://postgr.es/m/20200603021448.GB89559@paquier.xyz
Previously this code assumed that all IndexScan nodes supported
mark/restore, which is not true since it depends on optional index AM
support functions. This could lead to errors about missing support
functions in rare edge cases of mergejoins with no sort keys, where an
unordered non-btree index scan was placed on the inner path without a
protecting Materialize node. (Normally, the fact that merge join
requires ordered input would avoid this error.)
Backpatch all the way since this bug is ancient.
Per report from Eugen Konkov on irc.
Discussion: https://postgr.es/m/87o8jn50be.fsf@news-spur.riddles.org.uk
Gen_fmgrtab.pl treated aggregate functions the same as other built-in
functions, which is wasteful because there is no real need to have
entries for them in the fmgr_builtins[] table. Suppressing those
entries saves about 3KB in the compiled table on my machine; which
is not a lot but it's not nothing either, considering that that
table is pretty "hot". The only outside code change needed is
that ExecInitWindowAgg() can't be allowed to call fmgr_info_cxt()
on a plain aggregate function. But that saves a few cycles anyway.
Having done that, the aggregate_dummy() function is unreferenced
and might as well be dropped. Using "aggregate_dummy" as the prosrc
value for an aggregate is now just a documentation convention not
something that matters. There was some discussion of using NULL
instead to save a few bytes in pg_proc, but we'd have to remove
prosrc's BKI_FORCE_NOT_NULL marking which doesn't seem a great idea.
Anyway, it's possible there's client-side code that expects to
see "aggregate_dummy" there, so I'm loath to change it without a
strong reason.
Discussion: https://postgr.es/m/533989.1604263665@sss.pgh.pa.us
If the planner erroneously puts a non-parallel-safe SubPlan into
a parallelized portion of the query tree, nodeSubplan.c will fail
in the worker processes because it finds a null in es_subplanstates,
which it's unable to cope with. It seems worth a test-and-elog to
make that an error case rather than a core dump case.
This probably should have been included in commit 16ebab688, which
was responsible for allowing nulls to appear in es_subplanstates
to begin with. So, back-patch to v10 where that came in.
Discussion: https://postgr.es/m/924226.1604422326@sss.pgh.pa.us
Previously, ExecInitModifyTable relied on ExecInitJunkFilter,
and thence ExecCleanTypeFromTL, to build the target descriptor from
the query tlist. While we just checked (in ExecCheckPlanOutput)
that the tlist produces compatible output, this is not a great
substitute for the relation's actual tuple descriptor that's
available from the relcache. For one thing, dropped columns will
not be correctly marked attisdropped; it's a bit surprising that
we've gotten away with that this long. But the real reason for
being concerned with this is that using the table's descriptor means
that the slot will have correct attrmissing data, allowing us to
revert the klugy fix of commit ba9f18abd. (This commit undoes
that one's changes in trigger.c, but keeps the new test case.)
Thus we can solve the bogus-trigger-tuple problem with fewer cycles
rather than more.
No back-patch, since this doesn't fix any additional bug, and it
seems somewhat more likely to have unforeseen side effects than
ba9f18abd's narrow fix.
Discussion: https://postgr.es/m/16644-5da7ef98a7ac4545@postgresql.org
Since commit 913bbd88d, check_sql_fn_retval() can either insert type
coercion steps in-line in the Query that produces the SQL function's
results, or generate a new top-level Query to perform the coercions,
if modifying the Query's output in-place wouldn't be safe. However,
it appears that the latter case has never actually worked, because
the code tried to inject the new Query back into the query list it was
passed ... which is not the list that will be used for later processing
when we execute the SQL function "normally" (without inlining it).
So we ended up with no coercion happening at run-time, leading to
wrong results or crashes depending on the datatypes involved.
While the regression tests look like they cover this area well enough,
through a huge bit of bad luck all the test cases that exercise the
separate-Query path were checking either inline-able cases (which
accidentally didn't have the bug) or cases that are no-ops at runtime
(e.g., varchar to text), so that the failure to perform the coercion
wasn't obvious. The fact that the cases that don't work weren't
allowed at all before v13 probably contributed to not noticing the
problem sooner, too.
To fix, get rid of the separate "flat" list of Query nodes and instead
pass the real two-level list that is going to be used later. I chose
to make the same change in check_sql_fn_statements(), although that has
no actual bug, just so that we don't need that data structure at all.
This is an API change, as evidenced by the adjustments needed to
callers outside functions.c. That's a bit scary to be doing in a
released branch, but so far as I can tell from a quick search,
there are no outside callers of these functions (and they are
sufficiently specific to our semantics for SQL-language functions that
it's not apparent why any extension would need to call them). In any
case, v13 already changed the API of check_sql_fn_retval() compared to
prior branches.
Per report from pinker. Back-patch to v13 where this code came in.
Discussion: https://postgr.es/m/1603050466566-0.post@n3.nabble.com
Store the tuple conversion map to convert a tuple from a child table's
format to the root format in a new ri_ChildToRootMap field in
ResultRelInfo. It is initialized if transition tuple capture for FOR
STATEMENT triggers or INSERT tuple routing on a partitioned table is
needed. Previously, ModifyTable kept the maps in the per-subplan
ModifyTableState->mt_per_subplan_tupconv_maps array, or when tuple
routing was used, in
ResultRelInfo->ri_Partitioninfo->pi_PartitionToRootMap. The new field
replaces both of those.
Now that the child-to-root tuple conversion map is always available in
ResultRelInfo (when needed), remove the TransitionCaptureState.tcs_map
field. The callers of Exec*Trigger() functions no longer need to set or
save it, which is much less confusing and bug-prone. Also, as a future
optimization, this will allow us to delay creating the map for a given
result relation until the relation is actually processed during
execution.
Author: Amit Langote
Discussion: https://www.postgresql.org/message-id/CA%2BHiwqHtCWLdK-LO%3DNEsvOdHx%2B7yv4mE_zYK0i3BH7dXb-wxog%40mail.gmail.com
When executing DDL on a partitioned table or on a table with inheritance
children, statement-level triggers must be fired against the table given
in the original statement. The code to look that up was a bit messy and
duplicative. Commit 501ed02cf6 added a helper function,
getASTriggerResultRelInfo() (later renamed to getTargetResultRelInfo())
for it, but for some reason it was only used when firing AFTER STATEMENT
triggers and the code to fire BEFORE STATEMENT triggers duplicated the
logic.
Determine the target relation in ExecInitModifyTable(), and set it always
in ModifyTableState. Code that used to call getTargetResultRelInfo() can
now use ModifyTableState->rootResultRelInfo directly.
Discussion: https://www.postgresql.org/message-id/CA%2BHiwqFViT47Zbr_ASBejiK7iDG8%3DQ1swQ-tjM6caRPQ67pT%3Dw%40mail.gmail.com
Maintaining 'es_result_relation_info' correctly at all times has become
cumbersome, especially with partitioning where each partition gets its
own result relation info. Having to set and reset it across arbitrary
operations has caused bugs in the past.
This changes all the places that used 'es_result_relation_info', to
receive the currently active ResultRelInfo via function parameters
instead.
Author: Amit Langote
Discussion: https://www.postgresql.org/message-id/CA%2BHiwqGEmiib8FLiHMhKB%2BCH5dRgHSLc5N5wnvc4kym%2BZYpQEQ%40mail.gmail.com
FDWs that can perform an UPDATE/DELETE remotely using the "direct modify"
set of APIs need to access the ResultRelInfo of the target table. That's
currently available in EState.es_result_relation_info, but the next
commit will remove that field.
This commit adds a new resultRelation field in ForeignScan, to store the
target relation's RT index, and the corresponding ResultRelInfo in
ForeignScanState. The FDW's PlanDirectModify callback is expected to set
'resultRelation' along with 'operation'. The core code doesn't need them
for anything, they are for the convenience of FDW's Begin- and
IterateDirectModify callbacks.
Authors: Amit Langote, Etsuro Fujita
Discussion: https://www.postgresql.org/message-id/CA%2BHiwqGEmiib8FLiHMhKB%2BCH5dRgHSLc5N5wnvc4kym%2BZYpQEQ%40mail.gmail.com
Instead of allocating all the ResultRelInfos upfront in one big array,
allocate them in ExecInitModifyTable(). es_result_relations is now an
array of ResultRelInfo pointers, rather than an array of structs, and it
is indexed by the RT index.
This simplifies things: we get rid of the separate concept of a "result
rel index", and don't need to set it in setrefs.c anymore. This also
allows follow-up optimizations (not included in this commit yet) to skip
initializing ResultRelInfos for target relations that were not needed at
runtime, and removal of the es_result_relation_info pointer.
The EState arrays of regular result rels and root result rels are merged
into one array. Similarly, the resultRelations and rootResultRelations
lists in PlannedStmt are merged into one. It's not actually clear to me
why they were kept separate in the first place, but now that the
es_result_relations array is indexed by RT index, it certainly seems
pointless.
The PlannedStmt->resultRelations list is now only needed for
ExecRelationIsTargetRelation(). One visible effect of this change is that
ExecRelationIsTargetRelation() will now return 'true' also for the
partition root, if a partitioned table is updated. That seems like a good
thing, although the function isn't used in core code, and I don't see any
reason for an FDW to call it on a partition root.
Author: Amit Langote
Discussion: https://www.postgresql.org/message-id/CA%2BHiwqGEmiib8FLiHMhKB%2BCH5dRgHSLc5N5wnvc4kym%2BZYpQEQ%40mail.gmail.com
Unlike for functions, OUT parameters for procedures are part of the
signature. Therefore, they have to be listed in pg_proc.proargtypes
as well as mentioned in ALTER PROCEDURE and DROP PROCEDURE.
Reviewed-by: Andrew Dunstan <andrew.dunstan@2ndquadrant.com>
Reviewed-by: Pavel Stehule <pavel.stehule@gmail.com>
Discussion: https://www.postgresql.org/message-id/flat/2b8490fe-51af-e671-c504-47359dc453c5@2ndquadrant.com
When commit bd3daddaf introduced AlternativeSubPlans, I had some
ambitions towards allowing the choice of subplan to change during
execution. That has not happened, or even been thought about, in the
ensuing twelve years; so it seems like a failed experiment. So let's
rip that out and resolve the choice of subplan at the end of planning
(in setrefs.c) rather than during executor startup. This has a number
of positive benefits:
* Removal of a few hundred lines of executor code, since
AlternativeSubPlans need no longer be supported there.
* Removal of executor-startup overhead (particularly, initialization
of subplans that won't be used).
* Removal of incidental costs of having a larger plan tree, such as
tree-scanning and copying costs in the plancache; not to mention
setrefs.c's own costs of processing the discarded subplans.
* EXPLAIN no longer has to print a weird (and undocumented)
representation of an AlternativeSubPlan choice; it sees only the
subplan actually used. This should mean less confusion for users.
* Since setrefs.c knows which subexpression of a plan node it's
working on at any instant, it's possible to adjust the estimated
number of executions of the subplan based on that. For example,
we should usually estimate more executions of a qual expression
than a targetlist expression. The implementation used here is
pretty simplistic, because we don't want to expend a lot of cycles
on the issue; but it's better than ignoring the point entirely,
as the executor had to.
That last point might possibly result in shifting the choice
between hashed and non-hashed EXISTS subplans in a few cases,
but in general this patch isn't meant to change planner choices.
Since we're doing the resolution so late, it's really impossible
to change any plan choices outside the AlternativeSubPlan itself.
Patch by me; thanks to David Rowley for review.
Discussion: https://postgr.es/m/1992952.1592785225@sss.pgh.pa.us
Since there is only one place that actually needs the partition check
expression, namely ExecPartitionCheck, it's better to fetch it from
the relcache there. In this way we will never fetch it at all if
the query never has use for it, and we still fetch it just once when
we do need it.
The reason for taking an interest in this is that if the relcache
doesn't already have the check expression cached, fetching it
requires obtaining AccessShareLock on the partition root. That
means that operations that look like they should only touch the
partition itself will also take a lock on the root. In particular
we observed that TRUNCATE on a partition may take a lock on the
partition's root, contributing to a deadlock situation in parallel
pg_restore.
As written, this patch does have a small cost, which is that we
are microscopically reducing efficiency for the case where a partition
has an empty check expression. ExecPartitionCheck will be called,
and will go through the motions of setting up and checking an empty
qual, where before it would not have been called at all. We could
avoid that by adding a separate boolean flag to track whether there
is a partition expression to test. However, this case only arises
for a default partition with no siblings, which surely is not an
interesting case in practice. Hence adding complexity for it
does not seem like a good trade-off.
Amit Langote, per a suggestion by me
Discussion: https://postgr.es/m/VI1PR03MB31670CA1BD9625C3A8C5DD05EB230@VI1PR03MB3167.eurprd03.prod.outlook.com
Previously, it was based on nBlocksAllocated to account for tapes with
open write buffers that may not have made it to the BufFile yet.
That was unnecessary, because callers do not need to get the number of
blocks while a tape has an open write buffer; and it also conflicted
with the preallocation logic added for HashAgg.
Reviewed-by: Peter Geoghegan
Discussion: https://postgr.es/m/ce5af05900fdbd0e9185747825a7423c48501964.camel@j-davis.com
Backpatch-through: 13
Partitioning tuple route code assumes that the partition chosen while
descending the partition hierarchy is always the correct one. This is
true except when the partition is the default partition and another
partition has been added concurrently: the partition constraint changes
and we don't recheck it. This can lead to tuples mistakenly being added
to the default partition that should have been rejected.
Fix by rechecking the default partition constraint while descending the
hierarchy.
An isolation test based on the reproduction steps described by Hao Wu
(with tweaks for extra coverage) is included.
Backpatch to 12, where this bug came in with 898e5e3290.
Reported by: Hao Wu <hawu@vmware.com>
Author: Amit Langote <amitlangote09@gmail.com>
Author: Álvaro Herrera <alvherre@alvh.no-ip.org>
Discussion: https://postgr.es/m/CA+HiwqFqBmcSSap4sFnCBUEL_VfOMmEKaQ3gwUhyfa4c7J_-nA@mail.gmail.com
Discussion: https://postgr.es/m/DM5PR0501MB3910E97A9EDFB4C775CF3D75A42F0@DM5PR0501MB3910.namprd05.prod.outlook.com
nodeSubplan.c expects that the testexpr for a hashable ANY SubPlan
has the form of one or more OpExprs whose LHS is an expression of the
outer query's, while the RHS is an expression over Params representing
output columns of the subquery. However, the planner only went as far
as verifying that the clauses were all binary OpExprs. This works
99.99% of the time, because the clauses have the right shape when
emitted by the parser --- but it's possible for function inlining to
break that, as reported by PegoraroF10. To fix, teach the planner
to check that the LHS and RHS contain the right things, or more
accurately don't contain the wrong things. Given that this has been
broken for years without anyone noticing, it seems sufficient to just
give up hashing when it happens, rather than go to the trouble of
commuting the clauses back again (which wouldn't necessarily work
anyway).
While poking at that, I also noticed that nodeSubplan.c had a baked-in
assumption that the number of hash clauses is identical to the number
of subquery output columns. Again, that's fine as far as parser output
goes, but it's not hard to break it via function inlining. There seems
little reason for that assumption though --- AFAICS, the only thing
it's buying us is not having to store the number of hash clauses
explicitly. Adding code to the planner to reject such cases would take
more code than getting nodeSubplan.c to cope, so I fixed it that way.
This has been broken for as long as we've had hashable SubPlans,
so back-patch to all supported branches.
Discussion: https://postgr.es/m/1549209182255-0.post@n3.nabble.com
Since we no longer require AccessExclusiveLock to add a partition,
the executor may see that a partitioned table has more partitions
than the planner saw. ExecCreatePartitionPruneState's code for
matching up the partition lists in such cases was faulty, and would
misbehave if the planner had successfully pruned any partitions from
the query. (Thus, trouble would occur only if a partition addition
happens concurrently with a query that uses both static and dynamic
partition pruning.) This led to an Assert failure in debug builds,
and probably to crashes or query misbehavior in production builds.
To repair the bug, just explicitly skip zeroes in the plan's
relid_map[] list. I also made some cosmetic changes to make the code
more readable (IMO anyway). Also, convert the cross-checking Assert
to a regular test-and-elog, since it's now apparent that this logic
is more fragile than one would like.
Currently, there's no way to repeatably exercise this code, except
with manual use of a debugger to stop the backend between planning
and execution. Hence, no test case in this patch. We oughta do
something about that testability gap, but that's for another day.
Amit Langote and Tom Lane, per report from Justin Pryzby. Oversight
in commit 898e5e329; backpatch to v12 where that appeared.
Discussion: https://postgr.es/m/20200802181131.GA27754@telsasoft.com
Add a GUC that acts as a multiplier on work_mem. It gets applied when
sizing executor node hash tables that were previously size constrained
using work_mem alone.
The new GUC can be used to preferentially give hash-based nodes more
memory than the generic work_mem limit. It is intended to enable admin
tuning of the executor's memory usage. Overall system throughput and
system responsiveness can be improved by giving hash-based executor
nodes more memory (especially over sort-based alternatives, which are
often much less sensitive to being memory constrained).
The default value for hash_mem_multiplier is 1.0, which is also the
minimum valid value. This means that hash-based nodes continue to apply
work_mem in the traditional way by default.
hash_mem_multiplier is generally useful. However, it is being added now
due to concerns about hash aggregate performance stability for users
that upgrade to Postgres 13 (which added disk-based hash aggregation in
commit 1f39bce0). While the old hash aggregate behavior risked
out-of-memory errors, it is nevertheless likely that many users actually
benefited. Hash agg's previous indifference to work_mem during query
execution was not just faster; it also accidentally made aggregation
resilient to grouping estimate problems (at least in cases where this
didn't create destabilizing memory pressure).
hash_mem_multiplier can provide a certain kind of continuity with the
behavior of Postgres 12 hash aggregates in cases where the planner
incorrectly estimates that all groups (plus related allocations) will
fit in work_mem/hash_mem. This seems necessary because hash-based
aggregation is usually much slower when only a small fraction of all
groups can fit. Even when it isn't possible to totally avoid hash
aggregates that spill, giving hash aggregation more memory will reliably
improve performance (the same cannot be said for external sort
operations, which appear to be almost unaffected by memory availability
provided it's at least possible to get a single merge pass).
The PostgreSQL 13 release notes should advise users that increasing
hash_mem_multiplier can help with performance regressions associated
with hash aggregation. That can be taken care of by a later commit.
Author: Peter Geoghegan
Reviewed-By: Álvaro Herrera, Jeff Davis
Discussion: https://postgr.es/m/20200625203629.7m6yvut7eqblgmfo@alap3.anarazel.de
Discussion: https://postgr.es/m/CAH2-WzmD%2Bi1pG6rc1%2BCjc4V6EaFJ_qSuKCCHVnH%3DoruqD-zqow%40mail.gmail.com
Backpatch: 13-, where disk-based hash aggregation was introduced.
Use HyperLogLog to estimate the group cardinality in a spilled
partition. This estimate is used to choose the number of partitions if
we recurse.
The previous behavior was to use the number of tuples in a spilled
partition as the estimate for the number of groups, which lead to
overpartitioning. That could cause the number of batches to be much
higher than expected (with each batch being very small), which made it
harder to interpret EXPLAIN ANALYZE results.
Reviewed-by: Peter Geoghegan
Discussion: https://postgr.es/m/a856635f9284bc36f7a77d02f47bbb6aaf7b59b3.camel@j-davis.com
Backpatch-through: 13
There were various unnecessary differences between Hash Agg's EXPLAIN
ANALYZE output and Hash Join's. Here we modify the Hash Agg output so
that it's better aligned to Hash Join's.
The following changes have been made:
1. Start batches counter at 1 instead of 0.
2. Always display the "Batches" property, even when we didn't spill to
disk.
3. Use the text "Batches" instead of "HashAgg Batches" for text format.
4. Use the text "Memory Usage" instead of "Peak Memory Usage" for text
format.
5. Include "Batches" before "Memory Usage" in both text and non-text
formats.
In passing also modify the "Planned Partitions" property so that we show
it regardless of if the value is 0 or not for non-text EXPLAIN formats.
This was pointed out by Justin Pryzby and probably should have been part
of 40efbf870.
Reviewed-by: Justin Pryzby, Jeff Davis
Discussion: https://postgr.es/m/CAApHDvrshRnA6C0VFnu7Fb9TVvgGo80PUMm5+2DiaS1gEkPvtw@mail.gmail.com
Backpatch-through: 13, where HashAgg batching was introduced
Commit 85c9d347 addressed a similar problem for Gather and Gather
Merge nodes but forgot to account for nodes above parallel nodes. This
still works for nodes above Gather node because we shut down the workers
for Gather node as soon as there are no more tuples. We can do a similar
thing for Gather Merge as well but it seems better to account for stats
during nodes shutdown after completing the execution.
Reported-by: Stéphane Lorek, Jehan-Guillaume de Rorthais
Author: Jehan-Guillaume de Rorthais <jgdr@dalibo.com>
Reviewed-by: Amit Kapila
Backpatch-through: 10, where it was introduced
Discussion: https://postgr.es/m/20200718160206.584532a2@firost
The term "hash_mem" will take on new significance when pending work to
add a new hash_mem_multiplier GUC is committed. Rename a local variable
that happens to have been called hash_mem now to avoid confusion.
This representation saves 8 bytes per tuple compared to HeapTuple, and
avoids the need to allocate, copy and free on the receiving side.
Gather can emit the returned MinimalTuple directly, but GatherMerge now
needs to make an explicit copy because it buffers multiple tuples at a
time. That should be no worse than before.
Reviewed-by: Soumyadeep Chakraborty <soumyadeep2007@gmail.com>
Discussion: https://postgr.es/m/CA%2BhUKG%2B8T_ggoUTAE-U%3DA%2BOcPc4%3DB0nPPHcSfffuQhvXXjML6w%40mail.gmail.com
Since 1f39bce02, HashAgg nodes have had the ability to spill to disk when
memory consumption exceeds work_mem. That commit added new properties to
EXPLAIN ANALYZE to show the maximum memory usage and disk usage, however,
it didn't quite go as far as showing that information for parallel
workers. Since workers may have experienced something very different from
the main process, we should show this information per worker, as is done
in Sort.
Reviewed-by: Justin Pryzby
Reviewed-by: Jeff Davis
Discussion: https://postgr.es/m/CAApHDvpEKbfZa18mM1TD7qV6PG+w97pwCWq5tVD0dX7e11gRJw@mail.gmail.com
Backpatch-through: 13, where the hashagg spilling code was added.
Convert buffile.c error handling to use ereport. This fixes cases where
I/O errors were indistinguishable from EOF or not reported. Also remove
"%m" from error messages where errno would be bogus. While we're
modifying those strings, add block numbers and short read byte counts
where appropriate.
Back-patch to all supported releases.
Reported-by: Amit Khandekar <amitdkhan.pg@gmail.com>
Reviewed-by: Melanie Plageman <melanieplageman@gmail.com>
Reviewed-by: Alvaro Herrera <alvherre@2ndquadrant.com>
Reviewed-by: Robert Haas <robertmhaas@gmail.com>
Reviewed-by: Ibrar Ahmed <ibrar.ahmad@gmail.com>
Reviewed-by: Michael Paquier <michael@paquier.xyz>
Discussion: https://postgr.es/m/CA%2BhUKGJE04G%3D8TLK0DLypT_27D9dR8F1RQgNp0jK6qR0tZGWOw%40mail.gmail.com
plpgsql has always executed the query given in a RETURN QUERY command
by opening it as a cursor and then fetching a few rows at a time,
which it turns around and dumps into the function's result tuplestore.
The point of this was to keep from blowing out memory with an oversized
SPITupleTable result (note that while a tuplestore can spill tuples
to disk, SPITupleTable cannot). However, it's rather inefficient, both
because of extra data copying and because of executor entry/exit
overhead. In recent versions, a new performance problem has emerged:
use of a cursor prevents use of a parallel plan for the executed query.
We can improve matters by skipping use of a cursor and having the
executor push result tuples directly into the function's result
tuplestore. However, a moderate amount of new infrastructure is needed
to make that idea work:
* We can use the existing tstoreReceiver.c DestReceiver code to funnel
executor output to the tuplestore, but it has to be extended to support
plpgsql's requirement for possibly applying a tuple conversion map.
* SPI needs to be extended to allow use of a caller-supplied
DestReceiver instead of its usual receiver that puts tuples into
a SPITupleTable. Two new API calls are needed to handle both the
RETURN QUERY and RETURN QUERY EXECUTE cases.
I also felt that I didn't want these new API calls to use the legacy
method of specifying query parameter values with "char" null flags
(the old ' '/'n' convention); rather they should accept ParamListInfo
objects containing the parameter type and value info. This required
a bit of additional new infrastructure since we didn't yet have any
parse analysis callback that would interpret $N parameter symbols
according to type data supplied in a ParamListInfo. There seems to be
no harm in letting makeParamList install that callback by default,
rather than leaving a new ParamListInfo's parserSetup hook as NULL.
(Indeed, as of HEAD, I couldn't find anyplace that was using the
parserSetup field at all; plpgsql was using parserSetupArg for its
own purposes, but parserSetup seemed to be write-only.)
We can actually get plpgsql out of the business of using legacy null
flags altogether, and using ParamListInfo instead of its ad-hoc
PreparedParamsData structure; but this requires inventing one more
SPI API call that can replace SPI_cursor_open_with_args. That seems
worth doing, though.
SPI_execute_with_args and SPI_cursor_open_with_args are now unused
anywhere in the core PG distribution. Perhaps someday we could
deprecate/remove them. But cleaning up the crufty bits of the SPI
API is a task for a different patch.
Per bug #16040 from Jeremy Smith. This is unfortunately too invasive to
consider back-patching. Patch by me; thanks to Hamid Akhtar for review.
Discussion: https://postgr.es/m/16040-eaacad11fecfb198@postgresql.org
Synchronize the event names for parallel hash join waits with other
event names, by getting rid of the slashes and dropping "-ing"
suffixes. Rename ClogGroupUpdate to XactGroupUpdate, to match the
new SLRU name. Move the ProcSignalBarrier event to the IPC category;
it doesn't belong under IO.
Also a bit more wordsmithing in the wait event documentation tables.
Discussion: https://postgr.es/m/4505.1589640417@sss.pgh.pa.us
Thomas Munro fixed a longstanding annoyance in pg_bsd_indent, that
it would misformat lines containing IsA() macros on the assumption
that the IsA() call should be treated like a cast. This improves
some other cases involving field/variable names that match typedefs,
too. The only places that get worse are a couple of uses of the
OpenSSL macro STACK_OF(); we'll gladly take that trade-off.
Discussion: https://postgr.es/m/20200114221814.GA19630@alvherre.pgsql
In a logical replication subscriber, a table using REPLICA IDENTITY FULL
which has a primary key would try to use the primary key's index
available to scan for a tuple, but an assertion only assumed as correct
the case of an index associated to REPLICA IDENTITY USING INDEX. This
commit corrects the assertion so as the use of a primary key index is a
valid case.
Reported-by: Dilip Kumar
Analyzed-by: Dilip Kumar
Author: Euler Taveira
Reviewed-by: Michael Paquier, Masahiko Sawada
Discussion: https://postgr.es/m/CAFiTN-u64S5bUiPL1q5kwpHNd0hRnf1OE-bzxNiOs5zo84i51w@mail.gmail.com
Backpatch-through: 10
Includes some manual cleanup of places that pgindent messed up,
most of which weren't per project style anyway.
Notably, it seems some people didn't absorb the style rules of
commit c9d297751, because there were a bunch of new occurrences
of function calls with a newline just after the left paren, all
with faulty expectations about how the rest of the call would get
indented.
The additional pain from level 4 is excessive for the gain.
Also revert all the source annotation changes to their original
wordings, to avoid back-patching pain.
Discussion: https://postgr.es/m/31166.1589378554@sss.pgh.pa.us
Use it at level 4, a bit more restrictive than the default level, and
tweak our commanding comments to FALLTHROUGH.
(However, leave zic.c alone, since it's external code; to avoid the
warnings that would appear there, change CFLAGS for that file in the
Makefile.)
Author: Julien Rouhaud <rjuju123@gmail.com>
Author: Álvaro Herrera <alvherre@alvh.no-ip.org>
Reviewed-by: Tom Lane <tgl@sss.pgh.pa.us>
Discussion: https://postgr.es/m/20200412081825.qyo5vwwco3fv4gdo@nol
Discussion: https://postgr.es/m/flat/E1fDenm-0000C8-IJ@gemulon.postgresql.org
ExecReScanIncrementalSort was resetting bounded=false, which means the
optimization would be disabled on all rescans. This happens because
ExecSetTupleBound is called before the rescan, not after it.
Author: James Coleman
Reviewed-by: Tomas Vondra
Discussion: https://postgr.es/m/20200414065336.GI1492@paquier.xyz
The executor flags were not handled entirely correctly, although the
bugs were mostly harmless and it was mostly comment inaccuracy. We don't
need to strip any of the flags for child nodes.
Incremental sort does not support backward scans of mark/restore, so
MARK/BACKWARDS flags should not be possible. So we simply ensure this
using an assert, and we don't bother removing them when initializing
the child node.
With REWIND it's a bit less clear - incremental sort does not support
REWIND, but there is no way to signal this - it's legal to just ignore
the flag. We however continue passing the flag to child nodes, because
they might be useful to leverage that.
Reported-by: Michael Paquier
Author: James Coleman
Reviewed-by: Tomas Vondra
Discussion: https://postgr.es/m/20200414065336.GI1492@paquier.xyz
In commit 33e05f89c5, we have added the option to display WAL usage
statistics in Explain and auto_explain. The display format used two spaces
between each field which is inconsistent with Buffer usage statistics which
is using one space between each field. Change the format to make WAL usage
statistics consistent with Buffer usage statistics.
This commit also changed the usage of "full page writes" to
"full page images" for WAL usage statistics to make it consistent with
other parts of code and docs.
Author: Julien Rouhaud, Amit Kapila
Reviewed-by: Justin Pryzby, Kyotaro Horiguchi and Amit Kapila
Discussion: https://postgr.es/m/CAB-hujrP8ZfUkvL5OYETipQwA=e3n7oqHFU=4ZLxWS_Cza3kQQ@mail.gmail.com
In a9c35cf85c I changed ExecMakeTableFunctionResult() to dynamically
allocate the FunctionCallInfo used to call the SRF. Unfortunately I
did not account for the fact that the surrounding memory context has
query lifetime, leading to a leak till the end of the query.
In most cases the leak is fairly inconsequential, but if the
FunctionScan is done many times in the query, the leak can add
up. This happens e.g. if the function scan is on the inner side of a
nested loop, due to a lateral join.
EXPLAIN SELECT sum(f) FROM generate_series(1, 100000000) g(i), generate_series(i, i+1) f;
quickly shows the leak.
Instead of explicitly freeing the FunctionCallInfo it seems better to
make sure all the per-set temporary state in
ExecMakeTableFunctionResult() is cleaned up wholesale. Currently
that's probably just the FunctionCallInfo allocation, but since
there's some initialization work, and since there's already an
appropriate context, this seems like a more robust approach.
Bug: #16112
Reported-By: Ben Cornett
Author: Andres Freund
Reviewed-By: Tom Lane
Discussion: https://postgr.es/m/16112-4448bbf55a404189%40postgresql.org
Backpatch: 12, a9c35cf85c
Working on commit 1c455078b led me to check through FunctionCallInvoke
call sites to see if every one was being honest about (a) making sure
that fcinfo.isnull is initially false, and (b) checking its state after
the call. Sure enough, I found some violations.
The main one is that finalize_partialaggregate re-used serialfn_fcinfo
without resetting isnull, even though it clearly intends to cater for
serialfns that return NULL. There would only be an issue with a
non-strict serialfn, since it's unlikely that a serialfn would return
NULL for non-null input. We have no non-strict serialfns in core, and
there may be none in the wild either, which would account for the lack
of complaints. Still, it's clearly wrong, so back-patch that fix to
9.6 where finalize_partialaggregate was introduced.
Also, arrayfuncs.c and rowtypes.c contained various callers that were
not bothering to check for result nulls. While what's being called is
a comparison or hash function that probably *shouldn't* return null,
that's a lousy excuse for not having any check at all. There are
existing places that just Assert(!fcinfo->isnull) in comparable
situations, so I added that to the places that were calling btree
comparison or hash support functions. In the places calling
boolean-returning equality functions, it's quite cheap to have them
treat isnull as FALSE, so make those places do that. Also remove some
"locfcinfo->isnull = false" assignments that are unnecessary given the
assumption that no previous call returned null. These changes seem like
mostly neatnik-ism or debugging support, so I didn't back-patch.
In some corner cases, this could also lead to corrupted values being
included in the tuple.
Users who are concerned that they are affected by this should first
upgrade and then perform a base backup of their database and restore onto
an off-line server. They should then query each table with generated
columns to ensure there are no rows where the generated expression does
not match a newly calculated version of the GENERATED ALWAYS expression.
If no crashes occur and no rows are returned then you're not affected.
Fixes bug #16369.
Reported-by: Cameron Ezell
Discussion: https://postgr.es/m/16369-5845a6f1bef59884@postgresql.org
Backpatch-through: 12 (where GENERATED ALWAYS columns were added.)
The placement of the fall-through comment in this code appears not to
work to suppress the warning in recent gcc. Move it to the bottom of
the case group, and add an assertion that we didn't get there through
some other code path. Also improve wording of nearby comments.
Julien Rouhaud, comment hacking by me
Discussion: https://postgr.es/m/CAOBaU_aLdPGU5wCpaowNLF-Q8328iR7mj1yJAhMOVsdLwY+sdg@mail.gmail.com
Specifically, remember lookup_type_cache() results instead of retrieving
them once per comparison. Under CLOBBER_CACHE_ALWAYS, this reduced
src/test/subscription/t/001_rep_changes.pl elapsed time by an order of
magnitude, which reduced check-world elapsed time by 9%.
Discussion: https://postgr.es/m/20200406085420.GC162712@rfd.leadboat.com
Before discarding the old hash table in ExecReScanHashJoin, capture
its statistics, ensuring that we report the maximum hashtable size
across repeated rescans of the hash input relation. We can repurpose
the existing code for reporting hashtable size in parallel workers
to help with this, making the patch pretty small. This also ensures
that if rescans happen within parallel workers, we get the correct
maximums across all instances.
Konstantin Knizhnik and Tom Lane, per diagnosis by Thomas Munro
of a trouble report from Alvaro Herrera.
Discussion: https://postgr.es/m/20200323165059.GA24950@alvherre.pgsql
ExecReScanHashJoin will destroy the join's hash table if it expects
that the inner relation will produce different rows on rescan.
Up to now it's not bothered to clear the additional pointer to that
hash table that exists in the child HashState node. However, it's
possible for the query to terminate without building a fresh hash
table (this happens if the outer relation is found to be empty
during the final rescan). So we can end with a dangling pointer
to a deleted hash table. That was harmless originally, but since
9.0 EXPLAIN ANALYZE has used that pointer to print hash table
statistics. In debug builds this reproducibly results in garbage
statistics. In non-debug builds there's frequently no ill effects,
but in principle one could get wrong EXPLAIN ANALYZE output, or
perhaps even a crash if free() has released the hashtable memory
back to the OS.
To fix, just make sure we clear the additional pointer when destroying
the hash table. In problematic cases, EXPLAIN ANALYZE will then print
no hashtable statistics (reverting to its pre-9.0 behavior). This isn't
ideal, but since the problem manifests only in unusual corner cases,
it's hard to justify taking any risks to do better in the back
branches. A follow-on patch will improve matters in HEAD.
Konstantin Knizhnik and Tom Lane, per diagnosis by Thomas Munro
of a trouble report from Alvaro Herrera.
Discussion: https://postgr.es/m/20200323165059.GA24950@alvherre.pgsql
Currently, we don't account for buffer usage incurred by parallel workers
for parallel create index. This commit allows each worker to record the
buffer usage stats and leader backend to accumulate that stats at the
end of the operation. This will allow pg_stat_statements to display
correct buffer usage stats for (parallel) create index command.
Reported-by: Julien Rouhaud
Author: Sawada Masahiko
Reviewed-by: Dilip Kumar, Julien Rouhaud and Amit Kapila
Backpatch-through: 11, where this was introduced
Discussion: https://postgr.es/m/20200328151721.GB12854@nol
2nd pass of modifying various places which obtain the next power
of 2 of a number and make them use the new functions added in
f0705bb62.
In passing, also modify num_combinations(). This can be implemented
using simple bitshifting rather than looping.
Reviewed-by: John Naylor
Discussion: https://postgr.es/m/20200114173553.GE32763%40fetter.org
If the memory context's maxBlockSize is too big, a single block
allocation can suddenly exceed work_mem. For Hash Aggregation, this
can mean spilling to disk too early or reporting a confusing memory
usage number for EXPLAN ANALYZE.
Introduce CreateWorkExprContext(), which is like CreateExprContext(),
except that it creates the AllocSet with a maxBlockSize that is
reasonable in proportion to work_mem.
Right now, CreateWorkExprContext() is only used by Hash Aggregation,
but it may be generally useful in the future.
Discussion: https://postgr.es/m/412a3fbf306f84d8d78c4009e11791867e62b87c.camel@j-davis.com
WITH TIES is an option to the FETCH FIRST N ROWS clause (the SQL
standard's spelling of LIMIT), where you additionally get rows that
compare equal to the last of those N rows by the columns in the
mandatory ORDER BY clause.
There was a proposal by Andrew Gierth to implement this functionality in
a more powerful way that would yield more features, but the other patch
had not been finished at this time, so we decided to use this one for
now in the spirit of incremental development.
Author: Surafel Temesgen <surafel3000@gmail.com>
Reviewed-by: Álvaro Herrera <alvherre@alvh.no-ip.org>
Reviewed-by: Tomas Vondra <tomas.vondra@2ndquadrant.com>
Discussion: https://postgr.es/m/CALAY4q9ky7rD_A4vf=FVQvCGngm3LOes-ky0J6euMrg=_Se+ag@mail.gmail.com
Discussion: https://postgr.es/m/87o8wvz253.fsf@news-spur.riddles.org.uk
Incremental Sort is an optimized variant of multikey sort for cases when
the input is already sorted by a prefix of the requested sort keys. For
example when the relation is already sorted by (key1, key2) and we need
to sort it by (key1, key2, key3) we can simply split the input rows into
groups having equal values in (key1, key2), and only sort/compare the
remaining column key3.
This has a number of benefits:
- Reduced memory consumption, because only a single group (determined by
values in the sorted prefix) needs to be kept in memory. This may also
eliminate the need to spill to disk.
- Lower startup cost, because Incremental Sort produce results after each
prefix group, which is beneficial for plans where startup cost matters
(like for example queries with LIMIT clause).
We consider both Sort and Incremental Sort, and decide based on costing.
The implemented algorithm operates in two different modes:
- Fetching a minimum number of tuples without check of equality on the
prefix keys, and sorting on all columns when safe.
- Fetching all tuples for a single prefix group and then sorting by
comparing only the remaining (non-prefix) keys.
We always start in the first mode, and employ a heuristic to switch into
the second mode if we believe it's beneficial - the goal is to minimize
the number of unnecessary comparions while keeping memory consumption
below work_mem.
This is a very old patch series. The idea was originally proposed by
Alexander Korotkov back in 2013, and then revived in 2017. In 2018 the
patch was taken over by James Coleman, who wrote and rewrote most of the
current code.
There were many reviewers/contributors since 2013 - I've done my best to
pick the most active ones, and listed them in this commit message.
Author: James Coleman, Alexander Korotkov
Reviewed-by: Tomas Vondra, Andreas Karlsson, Marti Raudsepp, Peter Geoghegan, Robert Haas, Thomas Munro, Antonin Houska, Andres Freund, Alexander Kuzmenkov
Discussion: https://postgr.es/m/CAPpHfdscOX5an71nHd8WSUH6GNOCf=V7wgDaTXdDd9=goN-gfA@mail.gmail.com
Discussion: https://postgr.es/m/CAPpHfds1waRZ=NOmueYq0sx1ZSCnt+5QJvizT8ndT2=etZEeAQ@mail.gmail.com
Mainly, this adds support code in logical/worker.c for applying
replicated operations whose target is a partitioned table to its
relevant partitions.
Author: Amit Langote <amitlangote09@gmail.com>
Reviewed-by: Rafia Sabih <rafia.pghackers@gmail.com>
Reviewed-by: Peter Eisentraut <peter.eisentraut@2ndquadrant.com>
Reviewed-by: Petr Jelinek <petr@2ndquadrant.com>
Discussion: https://www.postgresql.org/message-id/flat/CA+HiwqH=Y85vRK3mOdjEkqFK+E=ST=eQiHdpj43L=_eJMOOznQ@mail.gmail.com
This allows gathering the WAL generation statistics for each statement
execution. The three statistics that we collect are the number of WAL
records, the number of full page writes and the amount of WAL bytes
generated.
This helps the users who have write-intensive workload to see the impact
of I/O due to WAL. This further enables us to see approximately what
percentage of overall WAL is due to full page writes.
In the future, we can extend this functionality to allow us to compute the
the exact amount of WAL data due to full page writes.
This patch in itself is just an infrastructure to compute WAL usage data.
The upcoming patches will expose this data via explain, auto_explain,
pg_stat_statements and verbose (auto)vacuum output.
Author: Kirill Bychik, Julien Rouhaud
Reviewed-by: Dilip Kumar, Fujii Masao and Amit Kapila
Discussion: https://postgr.es/m/CAB-hujrP8ZfUkvL5OYETipQwA=e3n7oqHFU=4ZLxWS_Cza3kQQ@mail.gmail.com
Don't try to be precise about it, just use a constant 16 bytes of
chunk overhead. Being smarter would require knowing the memory context
where the chunk will be allocated, which is not known by all callers.
Discussion: https://postgr.es/m/20200325220936.il3ni2fj2j2b45y5@alap3.anarazel.de
This commit adds query_string argument into the planner-related functions
and hook and allows us to pass the query string to them.
Currently there is no user of the query string passed. But the upcoming patch
for the planning counters will add the planning hook function into
pg_stat_statements and the function will need the query string. So this change
will be necessary for that patch.
Also this change is useful for some extensions that want to use the query
string in their planner hook function.
Author: Pascal Legrand, Julien Rouhaud
Reviewed-by: Yoshikazu Imai, Tom Lane, Fujii Masao
Discussion: https://postgr.es/m/CAOBaU_bU1m3_XF5qKYtSj1ua4dxd=FWDyh2SH4rSJAUUfsGmAQ@mail.gmail.com
Discussion: https://postgr.es/m/1583789487074-0.post@n3.nabble.com
Previously pg_stat_statements calculated the difference of buffer counters
by its own code even while BufferUsageAccumDiff() had the same code.
This commit expose BufferUsageAccumDiff() and makes pg_stat_statements
use it for the calculation, in order to simply the code.
This change also would be useful for the upcoming patch for the planning
counters in pg_stat_statements because the patch will add one more code
for the calculation of difference of buffer counters and that can easily be
done by using BufferUsageAccumDiff().
Author: Julien Rouhaud
Reviewed-by: Fujii Masao
Discussion: https://postgr.es/m/bdfee4e0-a304-2498-8da5-3cb52c0a193e@oss.nttdata.com
This reverts the parts of commit 17a28b0364
that changed ereport's auxiliary functions from returning dummy integer
values to returning void. It turns out that a minority of compilers
complain (not entirely unreasonably) about constructs such as
(condition) ? errdetail(...) : 0
if errdetail() returns void rather than int. We could update those
call sites to say "(void) 0" perhaps, but the expectation for this
patch set was that ereport callers would not have to change anything.
And this aspect of the patch set was already the most invasive and
least compelling part of it, so let's just drop it.
Per buildfarm.
Discussion: https://postgr.es/m/CA+fd4k6N8EjNvZpM8nme+y+05mz-SM8Z_BgkixzkA34R+ej0Kw@mail.gmail.com
Change all the auxiliary error-reporting routines to return void,
now that we no longer need to pretend they are passing something
useful to errfinish(). While this probably doesn't save anything
significant at the machine-code level, it allows detection of some
additional types of mistakes.
Pass the error location details (__FILE__, __LINE__, PG_FUNCNAME_MACRO)
to errfinish not errstart. This shaves a few cycles off the case where
errstart decides we're not going to emit anything.
Re-implement elog() as a trivial wrapper around ereport(), removing
the separate support infrastructure it used to have. Aside from
getting rid of some now-surplus code, this means that elog() now
really does have exactly the same semantics as ereport(), in particular
that it can skip evaluation work if the message is not to be emitted.
Andres Freund and Tom Lane
Discussion: https://postgr.es/m/CA+fd4k6N8EjNvZpM8nme+y+05mz-SM8Z_BgkixzkA34R+ej0Kw@mail.gmail.com
All errors of SQLSTATE class 23 should include the name of an object
associated with the error in separate fields of the error report message.
We do this so that applications need not try to extract them from the
possibly-localized human-readable text of the message.
Reported-by: Chris Bandy
Author: Chris Bandy
Reviewed-by: Amit Kapila and Amit Langote
Discussion: https://postgr.es/m/0aa113a3-3c7f-db48-bcd8-f9290b2269ae@gmail.com
While performing hash aggregation, track memory usage when adding new
groups to a hash table. If the memory usage exceeds work_mem, enter
"spill mode".
In spill mode, new groups are not created in the hash table(s), but
existing groups continue to be advanced if input tuples match. Tuples
that would cause a new group to be created are instead spilled to a
logical tape to be processed later.
The tuples are spilled in a partitioned fashion. When all tuples from
the outer plan are processed (either by advancing the group or
spilling the tuple), finalize and emit the groups from the hash
table. Then, create new batches of work from the spilled partitions,
and select one of the saved batches and process it (possibly spilling
recursively).
Author: Jeff Davis
Reviewed-by: Tomas Vondra, Adam Lee, Justin Pryzby, Taylor Vesely, Melanie Plageman
Discussion: https://postgr.es/m/507ac540ec7c20136364b5272acbcd4574aa76ef.camel@j-davis.com
It is possible to reach check_sql_fn_retval() with an unresolved
polymorphic rettype, resulting in an assertion failure as demonstrated
by one of the added test cases. However, the code following that
throws what seems an acceptable error message, so just remove the
Assert and adjust commentary.
While here, I thought it'd be a good idea to provide some parallel
tests of SQL-function and PL/pgSQL-function polymorphism behavior.
Some of these cases are perhaps duplicative of tests elsewhere,
but we hadn't any organized coverage of the topic AFAICS.
Although that assertion's been wrong all along, it won't have any
effect in production builds, so I'm not bothering to back-patch.
Discussion: https://postgr.es/m/21569.1584314271@sss.pgh.pa.us
The effective_io_concurrency GUC and equivalent tablespace option were
previously passed through a formula based on a theory about RAID
spindles and probabilities, to arrive at the number of pages to prefetch
in bitmap heap scans. Tomas Vondra, Andres Freund and others argued
that it was anachronistic and hard to justify, and commit 558a9165e0
already started down the path of bypassing it in new code. We agreed to
drop that logic and use the value directly.
For the default setting of 1, there is no change in effect. Higher
settings can be converted from the old meaning to the new with:
select round(sum(OLD / n::float)) from generate_series(1, OLD) s(n);
We might want to consider renaming the GUC before the next release given
the change in meaning, but it's not clear that many users had set it
very carefully anyway. That decision is deferred for now.
Discussion: https://postgr.es/m/CA%2BhUKGJUw08dPs_3EUcdO6M90GnjofPYrWp4YSLaBkgYwS-AqA%40mail.gmail.com
The need for this was removed by
8b9e9644dc.
A number of files now need to include utils/acl.h or
parser/parse_node.h explicitly where they previously got it indirectly
somehow.
Since parser/parse_node.h already includes nodes/parsenodes.h, the
latter is then removed where the former was added. Also, remove
nodes/pg_list.h from objectaddress.h, since that's included via
nodes/parsenodes.h.
Reviewed-by: Tom Lane <tgl@sss.pgh.pa.us>
Reviewed-by: Alvaro Herrera <alvherre@2ndquadrant.com>
Discussion: https://www.postgresql.org/message-id/flat/7601e258-26b2-8481-36d0-dc9dca6f28f1%402ndquadrant.com
Optionally push a step to check for a NULL pointer to the pergroup
state.
This will be important for disk-based hash aggregation in combination
with grouping sets. When memory limits are reached, a given tuple may
find its per-group state for some grouping sets but not others. For
the former, it advances the per-group state as normal; for the latter,
it skips evaluation and the calling code will have to spill the tuple
and reprocess it in a later batch.
Add the NULL check as a separate expression step because in some
common cases it's not needed.
Discussion: https://postgr.es/m/20200221202212.ssb2qpmdgrnx52sj%40alap3.anarazel.de
Our usual practice for "poor man's enum" catalog columns is to define
macros for the possible values and use those, not literal constants,
in C code. But for some reason lost in the mists of time, this was
never done for typalign/attalign or typstorage/attstorage. It's never
too late to make it better though, so let's do that.
The reason I got interested in this right now is the need to duplicate
some uses of the TYPSTORAGE constants in an upcoming ALTER TYPE patch.
But in general, this sort of change aids greppability and readability,
so it's a good idea even without any specific motivation.
I may have missed a few places that could be converted, and it's even
more likely that pending patches will re-introduce some hard-coded
references. But that's not fatal --- there's no expectation that
we'd actually change any of these values. We can clean up stragglers
over time.
Discussion: https://postgr.es/m/16457.1583189537@sss.pgh.pa.us
The backend was using strings to represent command tags and doing string
comparisons in multiple places, but that's slow and unhelpful. Create a
new command list with a supporting structure to use instead; this is
stored in a tag-list-file that can be tailored to specific purposes with
a caller-definable C macro, similar to what we do for WAL resource
managers. The first first such uses are a new CommandTag enum and a
CommandTagBehavior struct.
Replace numerous occurrences of char *completionTag with a
QueryCompletion struct so that the code no longer stores information
about completed queries in a cstring. Only at the last moment, in
EndCommand(), does this get converted to a string.
EventTriggerCacheItem no longer holds an array of palloc’d tag strings
in sorted order, but rather just a Bitmapset over the CommandTags.
Author: Mark Dilger, with unsolicited help from Álvaro Herrera
Reviewed-by: John Naylor, Tom Lane
Discussion: https://postgr.es/m/981A9DB4-3F0C-4DA5-88AD-CB9CFF4D6CAD@enterprisedb.com
Commit 356687bd8 omitted to remove leftover code for destroying
a hashed subplan's hash tables, with the result that the tables
were always rebuilt not reused; this leads to severe memory
leakage if a hashed subplan is re-executed enough times.
Moreover, the code for reusing the hashnulls table had a typo
that would have made it do the wrong thing if it were reached.
Looking at the code coverage report shows severe under-coverage
of the potential callers of ResetTupleHashTable, so add some test
cases that exercise them.
Andreas Karlsson and Tom Lane, per reports from Ranier Vilela
and Justin Pryzby.
Backpatch to v11, as the faulty commit was.
Discussion: https://postgr.es/m/edb62547-c453-c35b-3ed6-a069e4d6b937@proxel.se
Discussion: https://postgr.es/m/CAEudQAo=DCebm1RXtig9OH+QivpS97sMkikt0A9qHmMUs+g6ZA@mail.gmail.com
Discussion: https://postgr.es/m/20200210032547.GA1412@telsasoft.com
This also involves renaming src/include/utils/hashutils.h, which
becomes src/include/common/hashfn.h. Perhaps an argument can be
made for keeping the hashutils.h name, but it seemed more
consistent to make it match the name of the file, and also more
descriptive of what is actually going on here.
Patch by me, reviewed by Suraj Kharage and Mark Dilger. Off-list
advice on how not to break the Windows build from Davinder Singh
and Amit Kapila.
Discussion: http://postgr.es/m/CA+TgmoaRiG4TXND8QuM6JXFRkM_1wL2ZNhzaUKsuec9-4yrkgw@mail.gmail.com
Do so by combining the various steps that are part of aggregate
transition function invocation into one larger step. As some of the
current steps are only necessary for some aggregates, have one variant
of the aggregate transition step for each possible combination.
To avoid further manual copies of code in the different transition
step implementations, move most of the code into helper functions
marked as "always inline".
The benefit of this change is an increase in performance when
aggregating lots of rows. This comes in part due to the reduced number
of indirect jumps due to the reduced number of steps, and in part by
reducing redundant setup code across steps. This mainly benefits
interpreted execution, but the code generated by JIT is also improved
a bit.
As a nice side-effect it also ends up making the code a bit simpler.
A small additional optimization is removing the need to set
aggstate->curaggcontext before calling ExecAggInitGroup, choosing to
instead passign curaggcontext as an argument. It was, in contrast to
other aggregate related functions, only needed to fetch a memory
context to copy the transition value into.
Author: Andres Freund
Discussion:
https://postgr.es/m/20191023163849.sosqbfs5yenocez3@alap3.anarazel.dehttps://postgr.es/m/5c371df7cee903e8cd4c685f90c6c72086d3a2dc.camel@j-davis.com
* Separate calculation of hash value from the lookup.
* Split build_hash_table() into two functions.
* Change lookup_hash_entry() to return AggStatePerGroup. That's all
the caller needed, anyway.
These changes are to support the upcoming Disk-based Hash Aggregation
work.
Discussion: https://postgr.es/m/31f5ab871a3ad5a1a91a7a797651f20e77ac7ce3.camel%40j-davis.com
When updating a table row with generated columns, only recompute those
generated columns whose base columns have changed in this update and
keep the rest unchanged. This can result in a significant performance
benefit. The required information was already kept in
RangeTblEntry.extraUpdatedCols; we just have to make use of it.
Reviewed-by: Pavel Stehule <pavel.stehule@gmail.com>
Discussion: https://www.postgresql.org/message-id/flat/b05e781a-fa16-6b52-6738-761181204567@2ndquadrant.com
Commit 4eaea3db introduced TupleHashTableHash(), but the signature
didn't match the other exposed functions. Separate it into internal
and external versions. The external version hides the details behind
an API more consistent with the other external functions, and the
internal version is still suitable for simplehash.
Commit 147e3722f7 changed Tid scan so that it calls table_beginscan()
and uses the scan option for seq scan. This change caused two issues.
(1) The change caused Tid scan to take a predicate lock on the entire
relation in serializable transaction even when relation-level
lock is not necessary. This could lead to an unexpected
serialization error.
(2) The change caused Tid scan to increment the number of seq_scan
in pg_stat_*_tables views even though it's not seq scan. This
could confuse the users.
This commit adds the scan option for Tid scan and makes Tid scan
use it, to avoid those issues.
Back-patch to v12, where the bug was introduced.
Author: Tatsuhito Kasahara
Reviewed-by: Kyotaro Horiguchi, Masahiko Sawada, Fujii Masao
Discussion: https://postgr.es/m/CAP0=ZVKy+gTbFmB6X_UW0pP3WaeJ-fkUWHoD-pExS=at3CY76g@mail.gmail.com
Expose two new entry points: one for only calculating the hash value
of a tuple, and another for looking up a hash entry when the hash
value is already known. This will be useful for disk-based Hash
Aggregation to avoid recomputing the hash value for the same tuple
after saving and restoring it from disk.
Discussion: https://postgr.es/m/37091115219dd522fd9ed67333ee8ed1b7e09443.camel%40j-davis.com
It's already tracked via ExprState->parent, so we don't need to also
include it in ExprEvalStep. When that code originally was written
ExprState->parent didn't exist, but it since has been introduced in
6719b238e8.
Author: Andres Freund
Discussion: https://postgr.es/m/20191023163849.sosqbfs5yenocez3@alap3.anarazel.de
This mostly consists of using C99 style for loops, moving variables
into narrower scopes, and a smattering of other minor improvements.
Done separately to make it easier to review patches with actual
functional changes.
Author: Andres Freund
Discussion: https://postgr.es/m/20191023163849.sosqbfs5yenocez3@alap3.anarazel.de
When replica identity is FULL (an admittedly unusual case), the loop
that searches for tuples in execReplication.c didn't stop scanning the
table when once a matching tuple was found. Add the missing 'break'.
Note slight behavior change: we now return the first matching tuple
rather than the last one. They are supposed to be indistinguishable
anyway, so this shouldn't matter.
Author: Konstantin Knizhnik
Discussion: https://postgr.es/m/379743f6-ae91-b866-f7a2-5624e6d2b0a4@postgrespro.ru
We used to strategically place newlines after some function call left
parentheses to make pgindent move the argument list a few chars to the
left, so that the whole line would fit under 80 chars. However,
pgindent no longer does that, so the newlines just made the code
vertically longer for no reason. Remove those newlines, and reflow some
of those lines for some extra naturality.
Reviewed-by: Michael Paquier, Tom Lane
Discussion: https://postgr.es/m/20200129200401.GA6303@alvherre.pgsql
EvalPlanQualStart() supposed that it could re-use the relsubs_rowmark
and relsubs_done arrays from a prior instantiation. But since they are
allocated in the es_query_cxt of the recheckestate, that's just wrong;
EvalPlanQualEnd() will blow away that storage. Therefore we were using
storage that could have been reallocated to something else, causing all
sorts of havoc.
I think this was modeled on the old code's handling of es_epqTupleSlot,
but since the code was anyway clearing the arrays at re-use, there's
clearly no expectation of importing any outside state. So it's just
a dubious savings of a couple of pallocs, which is negligible compared
to setting up a new planstate tree. Therefore, just allocate the
arrays always. (I moved the allocations slightly for readability.)
In principle this bug could cause a problem whenever EPQ rechecks are
needed in more than one target table of a ModifyTable plan node.
In practice it seems not quite so easy to trigger as that; I couldn't
readily duplicate a crash with a partitioned target table, for instance.
That's probably down to incidental choices about when to free or
reallocate stuff. The added isolation test case does seem to reliably
show an assertion failure, though.
Per report from Oleksii Kliukin. Back-patch to v12 where the bug was
introduced (evidently by commit 3fb307bc4).
Discussion: https://postgr.es/m/EEF05F66-2871-4786-992B-5F45C92FEE2E@hintbits.com
This gives more information to the user about the error and it makes such
messages consistent with the other similar messages in the code.
Reported-by: Simon Riggs
Author: Mahendra Singh and Simon Riggs
Reviewed-by: Beena Emerson and Amit Kapila
Discussion: https://postgr.es/m/CANP8+j+7YUvQvGxTrCiw77R23enMJ7DFmyA3buR+fa2pKs4XhA@mail.gmail.com
Currently, Parallel Hash Join cannot be used for full/right joins,
so there is no point in setting the match flag. It turns out that
the cache coherence traffic generated by those writes slows down
large systems running many-core joins, so let's stop doing that.
In future, if we need to use match bits in parallel joins, we might
want to consider setting them only if not already set.
Back-patch to 11, where Parallel Hash Join arrived.
Reported-by: Deng, Gang
Discussion: https://postgr.es/m/0F44E799048C4849BAE4B91012DB910462E9897A%40SHSMSX103.ccr.corp.intel.com
The code checking whether an aggregate transition value needs to be
reparented into the current context has always only compared the
transition return value with the previous transition value by datum,
i.e. without regard for NULLness. This normally works, because when
the transition function returns NULL (via fcinfo->isnull), it'll
return a value that won't be the same as its input value.
But there's no hard requirement that that's the case. And it turns
out, it's possible to hit this case (see discussion or reproducers),
leading to a non-null transition value not being reparented, followed
by a crash caused by that.
Instead of adding another comparison of NULLness, instead have
ExecAggTransReparent() ensure that pergroup->transValue ends up as 0
when the new transition value is NULL. That avoids having to add an
additional branch to the much more common cases of the transition
function returning the old transition value (which is a pointer in
this case), and when the new value is different, but not NULL.
In branches since 69c3936a14, also deduplicate the reparenting code
between the expression evaluation based transitions, and the path for
ordered aggregates.
Reported-By: Teodor Sigaev, Nikita Glukhov
Author: Andres Freund
Discussion: https://postgr.es/m/bd34e930-cfec-ea9b-3827-a8bc50891393@sigaev.ru
Backpatch: 9.4-, this issue has existed since at least 7.4
Commit 9b63c13f0 turns out to have been fundamentally misguided:
the parent node's subPlan list is by no means the only way in which
a child SubPlan node can be hooked into the outer execution state.
As shown in bug #16213 from Matt Jibson, we can also get short-lived
tuple table slots added to the outer es_tupleTable list. At this point
I have little faith that there aren't other possible connections as
well; the long time it took to notice this problem shows that this
isn't a heavily-exercised situation.
Therefore, revert that fix, returning to the coding that passed a
NULL parent plan pointer down to the transiently-built subexpressions.
That gives us a pretty good guarantee that they won't hook into the
outer executor state in any way. But then we need some other solution
to make SubPlans work. Adopt the solution speculated about in the
previous commit's log message: do expression initialization at plan
startup for just those VALUES rows containing SubPlans, abandoning the
goal of reclaiming memory intra-query for those rows. In practice it
seems unlikely that queries containing a vast number of VALUES rows
would be using SubPlans in them, so this should not give up much.
(BTW, this test case also refutes my claim in connection with the prior
commit that the issue only arises with use of LATERAL. That was just
wrong: some variants of SubLink always produce SubPlans.)
As with previous patch, back-patch to all supported branches.
Discussion: https://postgr.es/m/16213-871ac3bc208ecf23@postgresql.org
Previously, check_xact_readonly() was responsible for determining
which types of queries could not be run in a read-only transaction,
standard_ProcessUtility() was responsibility for prohibiting things
which were allowed in read only transactions but not in recovery, and
utility commands were basically prohibited in bulk in parallel mode by
calls to CommandIsReadOnly() in functions.c and spi.c. This situation
was confusing and error-prone. Accordingly, move all the checks to a
new function ClassifyUtilityCommandAsReadOnly(), which determines the
degree to which a given statement is read only.
In the old code, check_xact_readonly() inadvertently failed to handle
several statement types that actually should have been prohibited,
specifically T_CreatePolicyStmt, T_AlterPolicyStmt, T_CreateAmStmt,
T_CreateStatsStmt, T_AlterStatsStmt, and T_AlterCollationStmt. As a
result, thes statements were erroneously allowed in read only
transactions, parallel queries, and standby operation. Generally, they
would fail anyway due to some lower-level error check, but we
shouldn't rely on that. In the new code structure, future omissions
of this type should cause ClassifyUtilityCommandAsReadOnly() to
complain about an unrecognized node type.
As a fringe benefit, this means we can allow certain types of utility
commands in parallel mode, where it's safe to do so. This allows
ALTER SYSTEM, CALL, DO, CHECKPOINT, COPY FROM, EXPLAIN, and SHOW.
It might be possible to allow additional commands with more work
and thought.
Along the way, document the thinking process behind the current set
of checks, as per discussion especially with Peter Eisentraut. There
is some interest in revising some of these rules, but that seems
like a job for another patch.
Patch by me, reviewed by Tom Lane, Stephen Frost, and Peter
Eisentraut.
Discussion: http://postgr.es/m/CA+TgmoZ_rLqJt5sYkvh+JpQnfX0Y+B2R+qfi820xNih6x-FQOQ@mail.gmail.com
A view with conditional INSTEAD rules and no unconditional INSTEAD
rules or INSTEAD OF triggers is not auto-updatable. Previously we
relied on a check in the executor to catch this, but that's
problematic since the planner may fail to properly handle such a query
and thus return a particularly unhelpful error to the user, before
reaching the executor check.
Instead, trap this in the rewriter and report the correct error there.
Doing so also allows us to include more useful error detail than the
executor check can provide. This doesn't change the existing behaviour
of updatable views; it merely ensures that useful error messages are
reported when a view isn't updatable.
Per report from Pengzhou Tang, though not adopting that suggested fix.
Back-patch to all supported branches.
Discussion: https://postgr.es/m/CAG4reAQn+4xB6xHJqWdtE0ve_WqJkdyCV4P=trYr4Kn8_3_PEA@mail.gmail.com
Use the parser's standard type coercion machinery to convert the
output column(s) of a SQL function's final SELECT or RETURNING
to the type(s) they should have according to the function's declared
result type. We'll allow any case where an assignment-level
coercion is available. Previously, we failed unless the required
coercion was a binary-compatible one (and the documentation ignored
this, falsely claiming that the types must match exactly).
Notably, the coercion now accounts for typmods, so that cases where
a SQL function is declared to return a composite type whose columns
are typmod-constrained now behave as one would expect. Arguably
this aspect is a bug fix, but the overall behavioral change here
seems too large to consider back-patching.
A nice side-effect is that functions can now be inlined in a
few cases where we previously failed to do so because of type
mismatches.
Discussion: https://postgr.es/m/18929.1574895430@sss.pgh.pa.us
This follows multiple complains from Peter Geoghegan, Andres Freund and
Alvaro Herrera that this issue ought to be dug more before actually
happening, if it happens.
Discussion: https://postgr.es/m/20191226144606.GA5659@alvherre.pgsql
Formerly the rd_partkey and rd_partdesc data structures were always
populated immediately when a relcache entry was built or rebuilt.
This patch changes things so that they are populated only when they
are first requested. (Hence, callers *must* now always use
RelationGetPartitionKey or RelationGetPartitionDesc; just fetching
the pointer directly is no longer acceptable.)
This seems to have some performance benefits, but the main reason to do
it is that it eliminates a recursive-reload failure that occurs if the
partkey or partdesc expressions contain any references to the relation's
rowtype (as discovered by Amit Langote). In retrospect, since loading
these data structures might result in execution of nearly-arbitrary code
via eval_const_expressions, it was a dumb idea to require that to happen
during relcache entry rebuild.
Also, fix things so that old copies of a relcache partition descriptor
will be dropped when the cache entry's refcount goes to zero. In the
previous coding it was possible for such copies to survive for the
lifetime of the session, as I'd complained of in a previous discussion.
(This management technique still isn't perfect, but it's better than
before.) Improve the commentary explaining how that works and why
it's safe to hand out direct pointers to these relcache substructures.
In passing, improve RelationBuildPartitionDesc by using the same
memory-context-parent-swap approach used by RelationBuildPartitionKey,
thereby making it less dependent on strong assumptions about what
partition_bounds_copy does. Avoid doing get_rel_relkind in the
critical section, too.
Patch by Amit Langote and Tom Lane; Robert Haas deserves some credit
for prior work in the area, too. Although this is a pre-existing
problem, no back-patch: the patch seems too invasive to be safe to
back-patch, and the bug it fixes is a corner case that seems
relatively unlikely to cause problems in the field.
Discussion: https://postgr.es/m/CA+HiwqFUzjfj9HEsJtYWcr1SgQ_=iCAvQ=O2Sx6aQxoDu4OiHw@mail.gmail.com
Discussion: https://postgr.es/m/CA+TgmoY3bRmGB6-DUnoVy5fJoreiBJ43rwMrQRCdPXuKt4Ykaw@mail.gmail.com
The following renaming is done so as source files related to index
access methods are more consistent with table access methods (the
original names used for index AMs ware too generic, and could be
confused as including features related to table AMs):
- amapi.h -> indexam.h.
- amapi.c -> indexamapi.c. Here we have an equivalent with
backend/access/table/tableamapi.c.
- amvalidate.c -> indexamvalidate.c.
- amvalidate.h -> indexamvalidate.h.
- genam.c -> indexgenam.c.
- genam.h -> indexgenam.h.
This has been discussed during the development of v12 when table AM was
worked on, but the renaming never happened.
Author: Michael Paquier
Reviewed-by: Fabien Coelho, Julien Rouhaud
Discussion: https://postgr.es/m/20191223053434.GF34339@paquier.xyz
Using \ is unnecessary and ugly, so remove that. While at it, stitch
the literals back into a single line: we've long discouraged splitting
error message literals even when they go past the 80 chars line limit,
to improve greppability.
Leave contrib/tablefunc alone.
Discussion: https://postgr.es/m/20191223195156.GA12271@alvherre.pgsql
Our algorithm for choosing batch numbers turned out not to work
effectively for multi-billion key inner relations. We would use
more hash bits than we have, and effectively concentrate all tuples
into a smaller number of batches than we intended. While ideally
we should switch to wider hashes, for now, change the algorithm to
one that effectively gives up bits from the bucket number when we
don't have enough bits. That means we'll finish up with longer
bucket chains than would be ideal, but that's better than having
batches that don't fit in work_mem and can't be divided.
Batch-patch to all supported releases.
Author: Thomas Munro
Reviewed-by: Tom Lane, thanks also to Tomas Vondra, Alvaro Herrera, Andres Freund for testing and discussion
Reported-by: James Coleman
Discussion: https://postgr.es/m/16104-dc11ed911f1ab9df%40postgresql.org
Tuple conversion support in tupconvert.c is able to convert rowtypes
between two relations, inner and outer, which are logically equivalent
but have a different ordering or even dropped columns (used mainly for
inheritance tree and partitions). This makes use of attribute mappings,
which are simple arrays made of AttrNumber elements with a length
matching the number of attributes of the outer relation. The length of
the attribute mapping has been treated as completely independent of the
mapping itself until now, making it easy to pass down an incorrect
mapping length.
This commit refactors the code related to attribute mappings and moves
it into an independent facility called attmap.c, extracted from
tupconvert.c. This merges the attribute mapping with its length,
avoiding to try to guess what is the length of a mapping to use as this
is computed once, when the map is built.
This will avoid mistakes like what has been fixed in dc816e58, which has
used an incorrect mapping length by matching it with the number of
attributes of an inner relation (a child partition) instead of an outer
relation (a partitioned table).
Author: Michael Paquier
Reviewed-by: Amit Langote
Discussion: https://postgr.es/m/20191121042556.GD153437@paquier.xyz
Previously, if the startup pruning logic proved that all child nodes
of an Append or MergeAppend could be pruned, we still kept one, just
to keep EXPLAIN from failing. The previous commit removed the
ruleutils.c limitation that required this kluge, so drop it. That
results in less-confusing EXPLAIN output, as per a complaint from
Yuzuko Hosoya.
David Rowley
Discussion: https://postgr.es/m/001001d4f44b$2a2cca50$7e865ef0$@lab.ntt.co.jp
This patch causes EXPLAIN to always assign a separate table alias to the
parent RTE of an append relation (inheritance set); before, such RTEs
were ignored if not actually scanned by the plan. Since the child RTEs
now always have that same alias to start with (cf. commit 55a1954da),
the net effect is that the parent RTE usually gets the alias used or
implied by the query text, and the children all get that alias with "_N"
appended. (The exception to "usually" is if there are duplicate aliases
in different subtrees of the original query; then some of those original
RTEs will also have "_N" appended.)
This results in more uniform output for partitioned-table plans than
we had before: the partitioned table itself gets the original alias,
and all child tables have aliases with "_N", rather than the previous
behavior where one of the children would get an alias without "_N".
The reason for giving the parent RTE an alias, even if it isn't scanned
by the plan, is that we now use the parent's alias to qualify Vars that
refer to an appendrel output column and appear above the Append or
MergeAppend that computes the appendrel. But below the append, Vars
refer to some one of the child relations, and are displayed that way.
This seems clearer than the old behavior where a Var that could carry
values from any child relation was displayed as if it referred to only
one of them.
While at it, change ruleutils.c so that the code paths used by EXPLAIN
deal in Plan trees not PlanState trees. This effectively reverts a
decision made in commit 1cc29fe7c, which seemed like a good idea at
the time to make ruleutils.c consistent with explain.c. However,
it's problematic because we'd really like to allow executor startup
pruning to remove all the children of an append node when possible,
leaving no child PlanState to resolve Vars against. (That's not done
here, but will be in the next patch.) This requires different handling
of subplans and initplans than before, but is otherwise a pretty
straightforward change.
Discussion: https://postgr.es/m/001001d4f44b$2a2cca50$7e865ef0$@lab.ntt.co.jp
Revert part of commit 19df1702f5.
Early shutdown was added by that commit so that we could collect
statistics from workers, but unfortunately, it interacted badly with
rescans. The problem is that we ended up destroying the parallel context
which is required for rescans. This leads to rescans of a Limit node over
a Gather node to produce unpredictable results as it tries to access
destroyed parallel context. By reverting the early shutdown code, we
might lose statistics in some cases of Limit over Gather [Merge], but that
will require further study to fix.
Reported-by: Jerry Sievers
Diagnosed-by: Thomas Munro
Author: Amit Kapila, testcase by Vignesh C
Backpatch-through: 9.6
Discussion: https://postgr.es/m/87ims2amh6.fsf@jsievers.enova.com
Call ExecShutdownNode() after ExecutePlan()'s loop, rather than at each
break. We had forgotten to do that in one case. The omission caused
intermittent "temporary file leak" warnings from multi-batch parallel
hash joins with a LIMIT clause.
Back-patch to 11. Though the problem exists in theory in earlier
parallel query releases, nothing really depended on it.
Author: Kyotaro Horiguchi
Reviewed-by: Thomas Munro, Amit Kapila
Discussion: https://postgr.es/m/20191111.212418.2222262873417235945.horikyota.ntt%40gmail.com
Avoid creating transiently-inconsistent slot states where possible,
by not setting TTS_FLAG_SHOULDFREE until after the slot actually has
a free'able tuple pointer, and by making sure that we reset tts_nvalid
and related derived state before we replace the tuple contents. This
would only matter if something were to examine the slot after we'd
suffered some kind of error (e.g. out of memory) while manipulating
the slot. We typically don't do that, so these changes might just be
cosmetic --- but even if so, it seems like good future-proofing.
Also remove some redundant Asserts, and add a couple for consistency.
Back-patch to v12 where all this code was rewritten.
Discussion: https://postgr.es/m/16095-c3ff2e5283b8dba5@postgresql.org
When maintaining or merging patches, one of the most common sources
for conflicts are the list of objects in makefiles. Especially when
the split across lines has been changed on both sides, which is
somewhat common due to attempting to stay below 80 columns, those
conflicts are unnecessarily laborious to resolve.
By splitting, and alphabetically sorting, OBJS style lines into one
object per line, conflicts should be less frequent, and easier to
resolve when they still occur.
Author: Andres Freund
Discussion: https://postgr.es/m/20191029200901.vww4idgcxv74cwes@alap3.anarazel.de
The old names for the attribute-detoasting functions names included
the word "heap," which seems outdated now that the heap is only one of
potentially many table access methods.
On the other hand, toast_insert_or_update and toast_delete are
heap-specific, so rename them by adding "heap_" as a prefix.
Not all of the work of making the TOAST system fully accessible to AMs
other than the heap is done yet, but there seems to be little harm in
getting this renaming out of the way now. Commit
8b94dab066 already divided up the
functions among various files partially according to whether it was
intended that they should be heap-specific or AM-agnostic, so this is
just clarifying the division contemplated by that commit.
Patch by me, reviewed and tested by Prabhat Sabu, Thomas Munro,
Andres Freund, and Álvaro Herrera.
Discussion: http://postgr.es/m/CA+TgmoZv-=2iWM4jcw5ZhJeL18HF96+W1yJeYrnGMYdkFFnEpQ@mail.gmail.com
That avoids unnecessary work during both interpreted execution, and
JIT compiled expression evaluation. Both benefit from fewer expression
steps needing be processed, and for interpreted execution there now is
a fastpath dedicated to just fetching a value from a virtual
slot. That's e.g. beneficial for hashjoins over nodes that perform
projections, as the hashed columns are currently fetched individually.
Author: Soumyadeep Chakraborty, Andres Freund
Discussion: https://postgr.es/m/CAE-ML+9OKSN71+mHtfMD-L24oDp8dGTfaVjDU6U+j+FNAW5kRQ@mail.gmail.com
In the course of 5567d12ce0, 356687bd8 and 317ffdfeaa, I changed
BuildTupleHashTable[Ext]'s call to ExecBuildGroupingEqual to not pass
in the parent node, but NULL. Which in turn prevents the tuple
equality comparator from being JIT compiled. While that fixes
bug #15486, it is not actually necessary after all of the above commits,
as we don't re-build the comparator when using the new
BuildTupleHashTableExt() interface (as the content of the hashtable
are reset, but the TupleHashTable itself is not).
Therefore re-allow jit compilation for callers that use
BuildTupleHashTableExt with a separate context for "metadata" and
content.
As in the previous commit, there's ongoing work to make this easier to
test to prevent such regressions in the future, but that
infrastructure is not going to be backpatchable.
The performance impact of not JIT compiling hashtable equality
comparators can be substantial e.g. for aggregation queries that
aggregate a lot of input rows to few output rows (when there are a lot
of output groups, there will be fewer comparisons).
Author: Andres Freund
Discussion: https://postgr.es/m/20190927072053.njf6prdl3vb7y7qb@alap3.anarazel.de
Backpatch: 11, just as 5567d12ce0
For many queries the fact that the tuple descriptor from the lower
node was not taken into account when determining whether the type of a
slot is fixed, lead to tuple deforming for such upper nodes not to be
JIT accelerated.
I broke this in 675af5c01e.
There is ongoing work to enable writing regression tests for related
behavior (including a patch that would have detected this
regression), by optionally showing such details in EXPLAIN. But as it
seems unlikely that that will be suitable for stable branches, just
merge the fix for now.
While it's fairly close to the 12 release window, the fact that 11
continues to perform JITed tuple deforming in these cases, that
there's still cases where we do so in 12, and the fact that the
performance regression can be sizable, weigh in favor of fixing it
now.
Author: Andres Freund
Discussion: https://postgr.es/m/20190927072053.njf6prdl3vb7y7qb@alap3.anarazel.de
Backpatch: 12-, where 675af5c01e was merged.
The code used the destination slot's natts where it intended to
use the source slot's natts. Adding an Assert shows that there
is no case in "make check-world" where these counts are different,
so maybe this is a harmless bug, but it's still a bug.
Takayuki Tsunakawa
Discussion: https://postgr.es/m/0A3221C70F24FB45833433255569204D1FD34C0E@G01JPEXMBYT05
In ad0bda5d24 I changed the EvalPlanQual machinery to store
substitution tuples in slot, instead of using plain HeapTuples. The
main motivation for that was that using HeapTuples will be inefficient
for future tableams. But it turns out that that conversion was buggy
for non-locking rowmarks - the wrong tuple descriptor was used to
create the slot.
As a secondary issue 5db6df0c0 changed ExecLockRows() to begin EPQ
earlier, to allow to fetch the locked rows directly into the EPQ
slots, instead of having to copy tuples around. Unfortunately, as Tom
complained, that forces some expensive initialization to happen
earlier.
As a third issue, the test coverage for EPQ was clearly insufficient.
Fixing the first issue is unfortunately not trivial: Non-locked row
marks were fetched at the start of EPQ, and we don't have the type
information for the rowmarks available at that point. While we could
change that, it's not easy. It might be worthwhile to change that at
some point, but to fix this bug, it seems better to delay fetching
non-locking rowmarks when they're actually needed, rather than
eagerly. They're referenced at most once, and in cases where EPQ
fails, might never be referenced. Fetching them when needed also
increases locality a bit.
To be able to fetch rowmarks during execution, rather than
initialization, we need to be able to access the active EPQState, as
that contains necessary data. To do so move EPQ related data from
EState to EPQState, and, only for EStates creates as part of EPQ,
reference the associated EPQState from EState.
To fix the second issue, change EPQ initialization to allow use of
EvalPlanQualSlot() to be used before EvalPlanQualBegin() (but
obviously still requiring EvalPlanQualInit() to have been done).
As these changes made struct EState harder to understand, e.g. by
adding multiple EStates, significantly reorder the members, and add a
lot more comments.
Also add a few more EPQ tests, including one that fails for the first
issue above. More is needed.
Reported-By: yi huang
Author: Andres Freund
Reviewed-By: Tom Lane
Discussion:
https://postgr.es/m/CAHU7rYZo_C4ULsAx_LAj8az9zqgrD8WDd4hTegDTMM1LMqrBsg@mail.gmail.comhttps://postgr.es/m/24530.1562686693@sss.pgh.pa.us
Backpatch: 12-, where the EPQ changes were introduced
detoast.c/h contain functions required to detoast a datum, partially
or completely, plus a few other utility functions for examining the
size of toasted datums.
toast_internals.c/h contain functions that are used internally to the
TOAST subsystem but which (mostly) do not need to be accessed from
outside.
heaptoast.c/h contains code that is intrinsically specific to the
heap AM, either because it operates on HeapTuples or is based on the
layout of a heap page.
detoast.c and toast_internals.c are placed in
src/backend/access/common rather than src/backend/access/heap. At
present, both files still have dependencies on the heap, but that will
be improved in a future commit.
Patch by me, reviewed and tested by Prabhat Sabu, Thomas Munro,
Andres Freund, and Álvaro Herrera.
Discussion: http://postgr.es/m/CA+TgmoZv-=2iWM4jcw5ZhJeL18HF96+W1yJeYrnGMYdkFFnEpQ@mail.gmail.com
The message was included as a parameter when this function was added in
dcb2bda9b7, but I don't think it has ever served any useful purpose.
Let's stop spreading it pointlessly.
Reviewed by Amit Langote and Peter Eisentraut.
Discussion: https://postgr.es/m/20190806224728.GA17233@alvherre.pgsql
For most uses of acl.h the details of how "Acl" internally looks like
are irrelevant. It might make sense to move a lot of the
implementation details into a separate header at a later point.
The main motivation of this change is to avoid including fmgr.h (via
array.h, which needs it for exposed structs) in a lot of files that
otherwise don't need it. A subsequent commit will remove the fmgr.h
include from a lot of files.
Directly include utils/array.h and utils/expandeddatum.h from the
files that need them, but previously included them indirectly, via
acl.h.
Author: Andres Freund
Discussion: https://postgr.es/m/20190803193733.g3l3x3o42uv4qj7l@alap3.anarazel.de
Now that list_nth is O(1), there's no good reason to maintain a
separate array of RTE pointers rather than indexing into
estate->es_range_table. Deleting the array doesn't save all that
much either; but just on cleanliness grounds, it's better not to
have duplicate representations of the identical information.
Discussion: https://postgr.es/m/14960.1565384592@sss.pgh.pa.us
In the wake of commit 1cff1b95a, the result of list_concat no longer
shares the ListCells of the second input. Therefore, we can replace
"list_concat(x, list_copy(y))" with just "list_concat(x, y)".
To improve call sites that were list_copy'ing the first argument,
or both arguments, invent "list_concat_copy()" which produces a new
list sharing no ListCells with either input. (This is a bit faster
than "list_concat(list_copy(x), y)" because it makes the result list
the right size to start with.)
In call sites that were not list_copy'ing the second argument, the new
semantics mean that we are usually leaking the second List's storage,
since typically there is no remaining pointer to it. We considered
inventing another list_copy variant that would list_free the second
input, but concluded that for most call sites it isn't worth worrying
about, given the relative compactness of the new List representation.
(Note that in cases where such leakage would happen, the old code
already leaked the second List's header; so we're only discussing
the size of the leak not whether there is one. I did adjust two or
three places that had been troubling to free that header so that
they manually free the whole second List.)
Patch by me; thanks to David Rowley for review.
Discussion: https://postgr.es/m/11587.1550975080@sss.pgh.pa.us
Commit bf6c614a2 rearranged the lookup of the comparison operators
needed in a hashed subplan, and in so doing, broke the cross-type
case: it caused the original LHS-vs-RHS operator to be used to compare
hash table entries too (which of course are all of the RHS type).
This leads to C functions being passed a Datum that is not of the
type they expect, with the usual hazards of crashes and unauthorized
server memory disclosure.
For the set of hashable cross-type operators present in v11 core
Postgres, this bug is nearly harmless on 64-bit machines, which
may explain why it escaped earlier detection. But it is a live
security hazard on 32-bit machines; and of course there may be
extensions that add more hashable cross-type operators, which
would increase the risk.
Reported by Andreas Seltenreich. Back-patch to v11 where the
problem came in.
Security: CVE-2019-10209
In 5f32b29c18 I changed the creation of HashState.hashkeys to
actually use HashState as the parent (instead of HashJoinState, which
was incorrect, as they were executed below HashState), to fix the
problem of hashkeys expressions otherwise relying on slot types
appropriate for HashJoinState, rather than HashState as would be
correct. That reliance was only introduced in 12, which is why it
previously worked to use HashJoinState as the parent (although I'd be
unsurprised if there were problematic cases).
Unfortunately that's not a sufficient solution, because before this
commit, the to-be-hashed expressions referenced inner/outer as
appropriate for the HashJoin, not Hash. That didn't have obvious bad
consequences, because the slots containing the tuples were put into
ecxt_innertuple when hashing a tuple for HashState (even though Hash
doesn't have an inner plan).
There are less common cases where this can cause visible problems
however (rather than just confusion when inspecting such executor
trees). E.g. "ERROR: bogus varno: 65000", when explaining queries
containing a HashJoin where the subsidiary Hash node's hash keys
reference a subplan. While normally hashkeys aren't displayed by
EXPLAIN, if one of those expressions references a subplan, that
subplan may be printed as part of the Hash node - which then failed
because an inner plan was referenced, and Hash doesn't have that.
It seems quite possible that there's other broken cases, too.
Fix the problem by properly splitting the expression for the HashJoin
and Hash nodes at plan time, and have them reference the proper
subsidiary node. While other workarounds are possible, fixing this
correctly seems easy enough. It was a pretty ugly hack to have
ExecInitHashJoin put the expression into the already initialized
HashState, in the first place.
I decided to not just split inner/outer hashkeys inside
make_hashjoin(), but also to separate out hashoperators and
hashcollations at plan time. Otherwise we would have ended up having
two very similar loops, one at plan time and the other during executor
startup. The work seems to more appropriately belong to plan time,
anyway.
Reported-By: Nikita Glukhov, Alexander Korotkov
Author: Andres Freund
Reviewed-By: Tom Lane, in an earlier version
Discussion: https://postgr.es/m/CAPpHfdvGVegF_TKKRiBrSmatJL2dR9uwFCuR+teQ_8tEXU8mxg@mail.gmail.com
Backpatch: 12-
These were introduced by pgindent due to fixe to broken
indentation (c.f. 8255c7a5ee). Previously the mis-indentation of
function prototypes was creatively used to reduce indentation in a few
places.
As that formatting only exists in master and REL_12_STABLE, it seems
better to fix it in both, rather than having some odd indentation in
v12 that somebody might copy for future patches or such.
Author: Andres Freund
Discussion: https://postgr.es/m/20190728013754.jwcbe5nfyt3533vx@alap3.anarazel.de
Backpatch: 12-
Since 15d8f8312 we assert that - and since 7ef04e4d2c, 4da597edf1
rely on - the slot type for an expression's
ecxt_{outer,inner,scan}tuple not changing, unless explicitly flagged
as such. That allows to either skip deforming (for a virtual tuple
slot) or optimize the code for JIT accelerated deforming
appropriately (for other known slot types).
This assumption was sometimes violated for grouping sets, when
nodeAgg.c internally uses tuplesorts, and the child node doesn't
return a TTSOpsMinimalTuple type slot. Detect that case, and flag that
the outer slot might not be "fixed".
It's probably worthwhile to optimize this further in the future, and
more granularly determine whether the slot is fixed. As we already
instantiate per-phase transition and equal expressions, we could
cheaply set the slot type appropriately for each phase. But that's a
separate change from this bugfix.
This commit does include a very minor optimization by avoiding to
create a slot for handling tuplesorts, if no such sorts are
performed. Previously we created that slot unnecessarily in the common
case of computing all grouping sets via hashing. The code looked too
confusing without that, as the conditions for needing a sort slot and
flagging that the slot type isn't fixed, are the same.
Reported-By: Ashutosh Sharma
Author: Andres Freund
Discussion: https://postgr.es/m/CAE9k0PmNaMD2oHTEAhRyxnxpaDaYkuBYkLa1dpOpn=RS0iS2AQ@mail.gmail.com
Backpatch: 12-, where the bug was introduced in 15d8f8312
After 277cb78983 ON CONFLICT ... SET ... RETURNING failed with
ERROR: virtual tuple table slot does not have system attributes
when taking the update path, as the slot used to insert into the
table (and then process RETURNING) was defined to be a virtual slot in
that commit. Virtual slots don't support system columns except for
tableoid and ctid, as the other system columns are AM dependent.
Fix that by using a slot of the table's type. Add tests for system
column accesses in ON CONFLICT ... RETURNING.
Reported-By: Roby, bisected to the relevant commit by Jeff Janes
Author: Andres Freund
Discussion: https://postgr.es/m/73436355-6432-49B1-92ED-1FE4F7E7E100@finefun.com.au
Backpatch: 12-, where the bug was introduced in 277cb78983
When we already know the length that we're going to append, then it
makes sense to use appendBinaryStringInfo instead of
appendStringInfoString so that the append can be performed with a simple
memcpy() using a known length rather than having to first perform a
strlen() call to obtain the length.
Discussion: https://postgr.es/m/CAKJS1f8+FRAM1s5+mAa3isajeEoAaicJ=4e0WzrH3tAusbbiMQ@mail.gmail.com
In nodeAppend.c and nodeMergeAppend.c there were some foreach loops which
looped over the list of subplans and only performed any work if the
subplan index was found in a Bitmapset. With the old linked list
implementation of List, this form made sense as accessing the Nth list
element was O(N). However, thanks to 1cff1b95a we now have array-based
lists, so accessing the Nth element has become O(1).
Here we make the most of the O(1) lookups and just loop over the set
members of the Bitmapset with bms_next_member(). This performs slightly
better when a small number of the list items are in the Bitmapset. Micro
benchmarks show that when the Bitmapset contains all or most of the list
items then the new code is ever so slightly slower. In practice, the cost
is so small that it's drowned out by various other things such as locking
the relations belonging to each subplan, etc.
The primary goal here is to leave better code examples around which benefit
better from the new list implementation.
Reviewed-by: Tom Lane
Discussion: https://postgr.es/m/CAKJS1f8ZcsLVgkF4wOfRyMYTcPgLFiUAOedFC+U2vK_aFZk-BA@mail.gmail.com
Now that commit fec0778c8 drew a clear line between public and private
fields in SPITupleTable, it seems pretty silly that the count of valid
tuples isn't on the public side of that line. The reason why not was
that there wasn't such a count. For reasons lost in the mists of time,
spi.c preferred to keep a count of remaining free entries in the array.
But that seems pretty pointless: it's unlike the way we handle similar
code everywhere else, and it involves extra subtractions that surely
outweigh having to do a comparison rather than test-for-zero to check
for array-full.
Hence, rearrange so that this code does the expansible array logic
the same as everywhere else, with a count of valid entries alongside
the allocated array length. And document the count as public.
I looked for core-code callers where it would make sense to start
relying on tuptable->numvals rather than the separate SPI_processed
variable. Right now there don't seem to be places where it'd be
a win to do so without more code restructuring than I care to
undertake today. In principle, though, having SPITupleTables be
fully self-contained should be helpful down the line.
Discussion: https://postgr.es/m/16852.1563395722@sss.pgh.pa.us
Formerly, lcons was about the same speed as lappend, but with the new
List implementation, that's not so; with a long List, data movement
imposes an O(N) cost on lcons and list_delete_first, but not lappend.
Hence, invent list_delete_last with semantics parallel to
list_delete_first (but O(1) cost), and change various places to use
lappend and list_delete_last where this can be done without much
violence to the code logic.
There are quite a few places that construct result lists using lcons not
lappend. Some have semantic rationales for that; I added comments about
it to a couple that didn't have them already. In many such places though,
I think the coding is that way only because back in the dark ages lcons
was faster than lappend. Hence, switch to lappend where this can be done
without causing semantic changes.
In ExecInitExprRec(), this results in aggregates and window functions that
are in the same plan node being executed in a different order than before.
Generally, the executions of such functions ought to be independent of
each other, so this shouldn't result in visibly different query results.
But if you push it, as one regression test case does, you can show that
the order is different. The new order seems saner; it's closer to
the order of the functions in the query text. And we never documented
or promised anything about this, anyway.
Also, in gistfinishsplit(), don't bother building a reverse-order list;
it's easy now to iterate backwards through the original list.
It'd be possible to go further towards removing uses of lcons and
list_delete_first, but it'd require more extensive logic changes,
and I'm not convinced it's worth it. Most of the remaining uses
deal with queues that probably never get long enough to be worth
sweating over. (Actually, I doubt that any of the changes in this
patch will have measurable performance effects either. But better
to have good examples than bad ones in the code base.)
Patch by me, thanks to David Rowley and Daniel Gustafsson for review.
Discussion: https://postgr.es/m/21272.1563318411@sss.pgh.pa.us
This is numbered take 7, and addresses a set of issues around:
- Fixes for typos and incorrect reference names.
- Removal of unneeded comments.
- Removal of unreferenced functions and structures.
- Fixes regarding variable name consistency.
Author: Alexander Lakhin
Discussion: https://postgr.es/m/10bfd4ac-3e7c-40ab-2b2e-355ed15495e8@gmail.com
Originally, Postgres Lists were a more or less exact reimplementation of
Lisp lists, which consist of chains of separately-allocated cons cells,
each having a value and a next-cell link. We'd hacked that once before
(commit d0b4399d8) to add a separate List header, but the data was still
in cons cells. That makes some operations -- notably list_nth() -- O(N),
and it's bulky because of the next-cell pointers and per-cell palloc
overhead, and it's very cache-unfriendly if the cons cells end up
scattered around rather than being adjacent.
In this rewrite, we still have List headers, but the data is in a
resizable array of values, with no next-cell links. Now we need at
most two palloc's per List, and often only one, since we can allocate
some values in the same palloc call as the List header. (Of course,
extending an existing List may require repalloc's to enlarge the array.
But this involves just O(log N) allocations not O(N).)
Of course this is not without downsides. The key difficulty is that
addition or deletion of a list entry may now cause other entries to
move, which it did not before.
For example, that breaks foreach() and sister macros, which historically
used a pointer to the current cons-cell as loop state. We can repair
those macros transparently by making their actual loop state be an
integer list index; the exposed "ListCell *" pointer is no longer state
carried across loop iterations, but is just a derived value. (In
practice, modern compilers can optimize things back to having just one
loop state value, at least for simple cases with inline loop bodies.)
In principle, this is a semantics change for cases where the loop body
inserts or deletes list entries ahead of the current loop index; but
I found no such cases in the Postgres code.
The change is not at all transparent for code that doesn't use foreach()
but chases lists "by hand" using lnext(). The largest share of such
code in the backend is in loops that were maintaining "prev" and "next"
variables in addition to the current-cell pointer, in order to delete
list cells efficiently using list_delete_cell(). However, we no longer
need a previous-cell pointer to delete a list cell efficiently. Keeping
a next-cell pointer doesn't work, as explained above, but we can improve
matters by changing such code to use a regular foreach() loop and then
using the new macro foreach_delete_current() to delete the current cell.
(This macro knows how to update the associated foreach loop's state so
that no cells will be missed in the traversal.)
There remains a nontrivial risk of code assuming that a ListCell *
pointer will remain good over an operation that could now move the list
contents. To help catch such errors, list.c can be compiled with a new
define symbol DEBUG_LIST_MEMORY_USAGE that forcibly moves list contents
whenever that could possibly happen. This makes list operations
significantly more expensive so it's not normally turned on (though it
is on by default if USE_VALGRIND is on).
There are two notable API differences from the previous code:
* lnext() now requires the List's header pointer in addition to the
current cell's address.
* list_delete_cell() no longer requires a previous-cell argument.
These changes are somewhat unfortunate, but on the other hand code using
either function needs inspection to see if it is assuming anything
it shouldn't, so it's not all bad.
Programmers should be aware of these significant performance changes:
* list_nth() and related functions are now O(1); so there's no
major access-speed difference between a list and an array.
* Inserting or deleting a list element now takes time proportional to
the distance to the end of the list, due to moving the array elements.
(However, it typically *doesn't* require palloc or pfree, so except in
long lists it's probably still faster than before.) Notably, lcons()
used to be about the same cost as lappend(), but that's no longer true
if the list is long. Code that uses lcons() and list_delete_first()
to maintain a stack might usefully be rewritten to push and pop at the
end of the list rather than the beginning.
* There are now list_insert_nth...() and list_delete_nth...() functions
that add or remove a list cell identified by index. These have the
data-movement penalty explained above, but there's no search penalty.
* list_concat() and variants now copy the second list's data into
storage belonging to the first list, so there is no longer any
sharing of cells between the input lists. The second argument is
now declared "const List *" to reflect that it isn't changed.
This patch just does the minimum needed to get the new implementation
in place and fix bugs exposed by the regression tests. As suggested
by the foregoing, there's a fair amount of followup work remaining to
do.
Also, the ENABLE_LIST_COMPAT macros are finally removed in this
commit. Code using those should have been gone a dozen years ago.
Patch by me; thanks to David Rowley, Jesper Pedersen, and others
for review.
Discussion: https://postgr.es/m/11587.1550975080@sss.pgh.pa.us
This addresses a couple of issues in the code:
- Typos and inconsistencies in comments and function declarations.
- Removal of unreferenced function declarations.
- Removal of unnecessary compile flags.
- A cleanup error in regressplans.sh.
Author: Alexander Lakhin
Discussion: https://postgr.es/m/0c991fdf-2670-1997-c027-772a420c4604@gmail.com
Etsuro Fujita
Tom Lane
David Rowley
Peter Eisentraut
Andrew Dunstan
Tomas Vondra
Michael Paquier
Alvaro Herrera
Bruce Momjian
Andres Freund
Amit Kapila
Thomas Munro
Heikki Linnakangas
Andrew Gierth
Peter Geoghegan
Jeff Davis
Alexander Korotkov
Noah Misch
Fujii Masao
Robert Haas
Dean Rasheed