Add missing infrastructure for this node type, notably in ruleutils.c where
its lack could demonstrably cause EXPLAIN to fail. Add outfuncs/readfuncs
support. (outfuncs support is useful today for debugging purposes. The
readfuncs support may never be needed, since at present it would only
matter for parallel query and NextValueExpr should never appear in a
parallelizable query; but it seems like a bad idea to have a primnode type
that isn't fully supported here.) Teach planner infrastructure that
NextValueExpr is a volatile, parallel-unsafe, non-leaky expression node
with cost cpu_operator_cost. Given its limited scope of usage, there
*might* be no live bug today from the lack of that knowledge, but it's
certainly going to bite us on the rear someday. Teach pg_stat_statements
about the new node type, too.
While at it, also teach cost_qual_eval() that MinMaxExpr, SQLValueFunction,
XmlExpr, and CoerceToDomain should be charged as cpu_operator_cost.
Failing to do this for SQLValueFunction was an oversight in my commit
0bb51aa96. The others are longer-standing oversights, but no time like the
present to fix them. (In principle, CoerceToDomain could have cost much
higher than this, but it doesn't presently seem worth trying to examine the
domain's constraints here.)
Modify execExprInterp.c to execute NextValueExpr as an out-of-line
function; it seems quite unlikely to me that it's worth insisting that
it be inlined in all expression eval methods. Besides, providing the
out-of-line function doesn't stop anyone from inlining if they want to.
Adjust some places where NextValueExpr support had been inserted with the
aid of a dartboard rather than keeping it in the same order as elsewhere.
Discussion: https://postgr.es/m/23862.1499981661@sss.pgh.pa.us
Don't move parenthesized lines to the left, even if that means they
flow past the right margin.
By default, BSD indent lines up statement continuation lines that are
within parentheses so that they start just to the right of the preceding
left parenthesis. However, traditionally, if that resulted in the
continuation line extending to the right of the desired right margin,
then indent would push it left just far enough to not overrun the margin,
if it could do so without making the continuation line start to the left of
the current statement indent. That makes for a weird mix of indentations
unless one has been completely rigid about never violating the 80-column
limit.
This behavior has been pretty universally panned by Postgres developers.
Hence, disable it with indent's new -lpl switch, so that parenthesized
lines are always lined up with the preceding left paren.
This patch is much less interesting than the first round of indent
changes, but also bulkier, so I thought it best to separate the effects.
Discussion: https://postgr.es/m/E1dAmxK-0006EE-1r@gemulon.postgresql.org
Discussion: https://postgr.es/m/30527.1495162840@sss.pgh.pa.us
Change pg_bsd_indent to follow upstream rules for placement of comments
to the right of code, and remove pgindent hack that caused comments
following #endif to not obey the general rule.
Commit e3860ffa4d wasn't actually using
the published version of pg_bsd_indent, but a hacked-up version that
tried to minimize the amount of movement of comments to the right of
code. The situation of interest is where such a comment has to be
moved to the right of its default placement at column 33 because there's
code there. BSD indent has always moved right in units of tab stops
in such cases --- but in the previous incarnation, indent was working
in 8-space tab stops, while now it knows we use 4-space tabs. So the
net result is that in about half the cases, such comments are placed
one tab stop left of before. This is better all around: it leaves
more room on the line for comment text, and it means that in such
cases the comment uniformly starts at the next 4-space tab stop after
the code, rather than sometimes one and sometimes two tabs after.
Also, ensure that comments following #endif are indented the same
as comments following other preprocessor commands such as #else.
That inconsistency turns out to have been self-inflicted damage
from a poorly-thought-through post-indent "fixup" in pgindent.
This patch is much less interesting than the first round of indent
changes, but also bulkier, so I thought it best to separate the effects.
Discussion: https://postgr.es/m/E1dAmxK-0006EE-1r@gemulon.postgresql.org
Discussion: https://postgr.es/m/30527.1495162840@sss.pgh.pa.us
The new indent version includes numerous fixes thanks to Piotr Stefaniak.
The main changes visible in this commit are:
* Nicer formatting of function-pointer declarations.
* No longer unexpectedly removes spaces in expressions using casts,
sizeof, or offsetof.
* No longer wants to add a space in "struct structname *varname", as
well as some similar cases for const- or volatile-qualified pointers.
* Declarations using PG_USED_FOR_ASSERTS_ONLY are formatted more nicely.
* Fixes bug where comments following declarations were sometimes placed
with no space separating them from the code.
* Fixes some odd decisions for comments following case labels.
* Fixes some cases where comments following code were indented to less
than the expected column 33.
On the less good side, it now tends to put more whitespace around typedef
names that are not listed in typedefs.list. This might encourage us to
put more effort into typedef name collection; it's not really a bug in
indent itself.
There are more changes coming after this round, having to do with comment
indentation and alignment of lines appearing within parentheses. I wanted
to limit the size of the diffs to something that could be reviewed without
one's eyes completely glazing over, so it seemed better to split up the
changes as much as practical.
Discussion: https://postgr.es/m/E1dAmxK-0006EE-1r@gemulon.postgresql.org
Discussion: https://postgr.es/m/30527.1495162840@sss.pgh.pa.us
David Rowley found that the "use the smallest per-column selectivity"
heuristic applied in some cases by get_foreign_key_join_selectivity()
was badly off if the FK columns are independent, producing estimates
much worse than we got before that code was added in 9.6.
One case where that heuristic was used was for LEFT and FULL outer joins
with the referenced rel on the outside of the join. But we should not
really need to special-case those here. eqjoinsel() never has had such a
special case; the correction is applied by calc_joinrel_size_estimate()
instead. Let's just estimate such cases like inner joins and rely on that
later adjustment. (I think there was something of a thinko here, in that
the comments seem to be thinking about the selectivity as defined for
semi/anti joins; but that shouldn't apply to left/full joins.) Add a
regression test exercising such a case to show that this is sane in
at least some cases.
The other case where we used that heuristic was for SEMI/ANTI outer joins,
either if the referenced rel was on the outside, or if it was on the inside
but was part of a join within the RHS. In either case, the FK doesn't give
us a lot of traction towards estimating the selectivity. To ensure that
we don't have regressions from what happened before 9.6, let's punt by
ignoring the FK in such cases and applying the traditional selectivity
calculation. (We might be able to improve on that later, but for now
I just want to be sure it's not worse than 9.5.)
Report and patch by David Rowley, simplified a bit by me. Back-patch
to 9.6 where this code was added.
Discussion: https://postgr.es/m/CAKJS1f8NO8oCDcxrteohG6O72uU1saEVT9qX=R8pENr5QWerXw@mail.gmail.com
In a CHECK clause, a null result means true, whereas in a WHERE clause
it means false. predtest.c provided different functions depending on
which set of semantics applied to the predicate being proved, but had
no option to control what a null meant in the clauses provided as
axioms. Add one.
Use that in the partitioning code when figuring out whether the
validation scan on a new partition can be skipped. Rip out the
old logic that attempted (not very successfully) to compensate
for the absence of the necessary support in predtest.c.
Ashutosh Bapat and Robert Haas, reviewed by Amit Langote and
incorporating feedback from Tom Lane.
Discussion: http://postgr.es/m/CAFjFpReT_kq_uwU_B8aWDxR7jNGE=P0iELycdq5oupi=xSQTOw@mail.gmail.com
I'd always assumed that backend/optimizer/geqo/'s remarkably poor
showing on code coverage metrics was because we weren't exercising
it much in the regression tests. But it turns out that a good chunk
of the problem is that there's a bunch of code that is physically
unreachable (because the calls to it are #ifdef'd out in geqo_main.c)
but is being built anyway. Making the called code have #if guards
similar to the calling code saves a couple of kilobytes of executable
size and should make the coverage numbers more reflective of reality.
It's arguable that we should just delete all the unused recombination
mechanisms altogether, but I didn't feel a need to go that far today.
When costing a nestloop with stop-at-first-inner-match semantics, and a
non-indexscan inner path, final_cost_nestloop() wants to charge the full
scan cost of the inner rel at least once, with additional scans charged
at inner_rescan_run_cost which might be less. However the logic for
doing this effectively assumed that outer_matched_rows is at least 1.
If it's zero, which is not unlikely for a small outer rel, we ended up
charging inner_run_cost plus N times inner_rescan_run_cost, as much as
double the correct charge for an outer rel with only one row that
we're betting won't be matched. (Unless the inner rel is materialized,
in which case it has very small inner_rescan_run_cost and the cost
is not so far off what it should have been.)
The upshot of this was that the planner had a tendency to select plans
that failed to make effective use of the stop-at-first-inner-match
semantics, and that might have Materialize nodes in them even when the
predicted number of executions of the Materialize subplan was only 1.
This was not so obvious before commit 9c7f5229a, because the case only
arose in connection with semi/anti joins where there's not freedom to
reverse the join order. But with the addition of unique-inner joins,
it could result in some fairly bad planning choices, as reported by
Teodor Sigaev. Indeed, some of the test cases added by that commit
have plans that look dubious on closer inspection, and are changed
by this patch.
Fix the logic to ensure that we don't charge for too many inner scans.
I chose to adjust it so that the full-freight scan cost is associated
with an unmatched outer row if possible, not a matched one, since that
seems like a better model of what would happen at runtime.
This is a longstanding bug, but given the lesser impact in back branches,
and the lack of field complaints, I won't risk a back-patch.
Discussion: https://postgr.es/m/CAKJS1f-LzkUsFxdJ_-Luy38orQ+AdEXM5o+vANR+-pHAWPSecg@mail.gmail.com
Otherwise, set_plan_refs() can get applied to the same list
multiple times through different references, leading to chaos.
Amit Langote, Dilip Kumar, and Robert Haas, reviewed by Ashutosh
Bapat. Original report by Sveinn Sveinsson.
Discussion: http://postgr.es/m/20170517141151.1435.79890@wrigleys.postgresql.org
Consistently refer to such an entry as a "statistics object", not just
"statistics" or "extended statistics". Previously we had a mismash of
terms, accompanied by utter confusion as to whether the term was
singular or plural. That's not only grating (at least to the ear of
a native English speaker) but could be outright misleading, eg in error
messages that seemed to be referring to multiple objects where only one
could be meant.
This commit fixes the code and a lot of comments (though I may have
missed a few). I also renamed two new SQL functions,
pg_get_statisticsextdef -> pg_get_statisticsobjdef
pg_statistic_ext_is_visible -> pg_statistics_obj_is_visible
to conform better with this terminology.
I have not touched the SGML docs other than fixing those function
names; the docs certainly need work but it seems like a separable task.
Discussion: https://postgr.es/m/22676.1494557205@sss.pgh.pa.us
If the inner relation can be proven unique, that is it can have no more
than one matching row for any row of the outer query, then we might as
well implement the semijoin as a plain inner join, allowing substantially
more freedom to the planner. This is a form of outer join strength
reduction, but it can't be implemented in reduce_outer_joins() because
we don't have enough info about the individual relations at that stage.
Instead do it much like remove_useless_joins(): once we've built base
relations, we can make another pass over the SpecialJoinInfo list and
get rid of any entries representing reducible semijoins.
This is essentially a followon to the inner-unique patch (commit 9c7f5229a)
and makes use of the proof machinery that that patch created. We need only
minor refactoring of innerrel_is_unique's API to support this usage.
Per performance complaint from Teodor Sigaev.
Discussion: https://postgr.es/m/f994fc98-389f-4a46-d1bc-c42e05cb43ed@sigaev.ru
The inner-unique patch (commit 9c7f5229a) supposed that if we're
considering a JOIN_UNIQUE_INNER join path, we can always set inner_unique
for the join, because the inner path produced by create_unique_path should
be unique relative to the outer relation. However, that's true only if
we're considering joining to the whole outer relation --- otherwise we may
be applying only some of the join quals, and so the inner path might be
non-unique from the perspective of this join. Adjust the test to only
believe that we can set inner_unique if we have the whole semijoin LHS on
the outer side.
There is more that can be done in this area, but this commit is only
intended to provide the minimal fix needed to get correct plans.
Per report from Teodor Sigaev. Thanks to David Rowley for preliminary
investigation.
Discussion: https://postgr.es/m/f994fc98-389f-4a46-d1bc-c42e05cb43ed@sigaev.ru
Even though no actual tuples are ever inserted into a partitioned
table (the actual tuples are in the partitions, not the partitioned
table itself), we still need to have a ResultRelInfo for the
partitioned table, or per-statement triggers won't get fired.
Amit Langote, per a report from Rajkumar Raghuwanshi. Reviewed by me.
Discussion: http://postgr.es/m/CAKcux6%3DwYospCRY2J4XEFuVy0L41S%3Dfic7rmkbsU-GXhhSbmBg%40mail.gmail.com
is_parallel_safe() supposed that the only relevant property of a SubPlan
was the parallel safety of the referenced subplan tree. This is wrong:
the testexpr or args subtrees might contain parallel-unsafe stuff, as
demonstrated by the test case added here. However, just recursing into the
subtrees fails in a different way: we'll typically find PARAM_EXEC Params
representing the subplan's output columns in the testexpr. The previous
coding supposed that any Param must be treated as parallel-restricted, so
that a naive attempt at fixing this disabled parallel pushdown of SubPlans
altogether. We must instead determine, for any visited Param, whether it
is one that would be computed by a surrounding SubPlan node; if so, it's
safe to push down along with the SubPlan node.
We might later be able to extend this logic to cope with Params used for
correlated subplans and other cases; but that's a task for v11 or beyond.
Tom Lane and Amit Kapila
Discussion: https://postgr.es/m/7064.1492022469@sss.pgh.pa.us
We were applying the use_physical_tlist optimization to all relation
scan plans, even those implemented by custom scan providers. However,
that's a bad idea for a couple of reasons. The custom provider might
be unable to provide columns that it hadn't expected to be asked for
(for example, the custom scan might depend on an index-only scan).
Even more to the point, there's no good reason to suppose that this
"optimization" is a win for a custom scan; whatever the custom provider
is doing is likely not based on simply returning physical heap tuples.
(As a counterexample, if the custom scan is an interface to a column store,
demanding all columns would be a huge loss.) If it is a win, the custom
provider could make that decision for itself and insert a suitable
pathtarget into the path, anyway.
Per discussion with Dmitry Ivanov. Back-patch to 9.5 where custom scan
support was introduced. The argument that the custom provider can adjust
the behavior by changing the pathtarget only applies to 9.6+, but on
balance it seems more likely that use_physical_tlist will hurt custom
scans than help them.
Discussion: https://postgr.es/m/e29ddd30-8ef9-4da5-a50b-2bb7b8c7198d@postgrespro.ru
We'd managed to avoid doing this so far, but it seems pretty obvious
that it would be forced on us some day, and this is much the cleanest
way of approaching the open problem that parallel-unsafe subplans are
being transmitted to parallel workers. Anyway there's no space cost
due to alignment considerations, and the time cost is pretty minimal
since we're just copying the flag from the corresponding Path node.
(At least in most cases ... some of the klugier spots in createplan.c
have to work a bit harder.)
In principle we could perhaps get rid of SubPlan.parallel_safe,
but I thought it better to keep that in case there are reasons to
consider a SubPlan unsafe even when its child plan is parallel-safe.
This patch doesn't actually do anything with the new flags, but
I thought I'd commit it separately anyway.
Note: although this touches outfuncs/readfuncs, there's no need for
a catversion bump because Plan trees aren't stored on disk.
Discussion: https://postgr.es/m/87tw5x4vcu.fsf@credativ.de
This extends the castNode() notation introduced by commit 5bcab1114 to
provide, in one step, extraction of a list cell's pointer and coercion to
a concrete node type. For example, "lfirst_node(Foo, lc)" is the same
as "castNode(Foo, lfirst(lc))". Almost half of the uses of castNode
that have appeared so far include a list extraction call, so this is
pretty widely useful, and it saves a few more keystrokes compared to the
old way.
As with the previous patch, back-patch the addition of these macros to
pg_list.h, so that the notation will be available when back-patching.
Patch by me, after an idea of Andrew Gierth's.
Discussion: https://postgr.es/m/14197.1491841216@sss.pgh.pa.us
Commit ac2b09508 was not terribly carefully reviewed. Band-aid it to
not fail on non-RestrictInfo input, per report from Andreas Seltenreich.
Also make it do something more reasonable with variable-free clauses,
and improve nearby comments.
Discussion: https://postgr.es/m/87inmf5rdx.fsf@credativ.de
If there can certainly be no more than one matching inner row for a given
outer row, then the executor can move on to the next outer row as soon as
it's found one match; there's no need to continue scanning the inner
relation for this outer row. This saves useless scanning in nestloop
and hash joins. In merge joins, it offers the opportunity to skip
mark/restore processing, because we know we have not advanced past the
first possible match for the next outer row.
Of course, the devil is in the details: the proof of uniqueness must
depend only on joinquals (not otherquals), and if we want to skip
mergejoin mark/restore then it must depend only on merge clauses.
To avoid adding more planning overhead than absolutely necessary,
the present patch errs in the conservative direction: there are cases
where inner_unique or skip_mark_restore processing could be used, but
it will not do so because it's not sure that the uniqueness proof
depended only on "safe" clauses. This could be improved later.
David Rowley, reviewed and rather heavily editorialized on by me
Discussion: https://postgr.es/m/CAApHDvqF6Sw-TK98bW48TdtFJ+3a7D2mFyZ7++=D-RyPsL76gw@mail.gmail.com
As reported by Sean Johnston in bug #14614, since 9.6 the planner can fail
due to trying to look up the referent of a Var with varno 0. This happens
because we generate such Vars in generate_append_tlist, for lack of any
better way to describe the output of a SetOp node. In typical situations
nothing really cares about that, but given nested set-operation queries
we will call estimate_num_groups on the output of the subquery, and that
wants to know what a Var actually refers to. That logic used to look at
subquery->targetList, but in commit 3fc6e2d7f I'd switched it to look at
subroot->processed_tlist, ie the actual output of the subquery plan not the
parser's idea of the result. It seemed like a good idea at the time :-(.
As a band-aid fix, change it back.
Really we ought to have an honest way of naming the outputs of SetOp steps,
which suggests that it'd be a good idea for the parser to emit an RTE
corresponding to each one. But that's a task for another day, and it
certainly wouldn't yield a back-patchable fix.
Report: https://postgr.es/m/20170407115808.25934.51866@wrigleys.postgresql.org
Follow on patch in the multi-variate statistics patch series.
CREATE STATISTICS s1 WITH (dependencies) ON (a, b) FROM t;
ANALYZE;
will collect dependency stats on (a, b) and then use the measured
dependency in subsequent query planning.
Commit 7b504eb282 added
CREATE STATISTICS with n-distinct coefficients. These are now
specified using the mutually exclusive option WITH (ndistinct).
Author: Tomas Vondra, David Rowley
Reviewed-by: Kyotaro HORIGUCHI, Álvaro Herrera, Dean Rasheed, Robert Haas
and many other comments and contributions
Discussion: https://postgr.es/m/56f40b20-c464-fad2-ff39-06b668fac47c@2ndquadrant.com
Currently, the only type of child relation is an "other member rel",
which is the child of a baserel, but in the future joins and even
upper relations may have child rels. To facilitate that, introduce
macros that test to test for particular RelOptKind values, and use
them in various places where they help to clarify the sense of a test.
(For example, a test may allow RELOPT_OTHER_MEMBER_REL either because
it intends to allow child rels, or because it intends to allow simple
rels.)
Also, remove find_childrel_top_parent, which will not work for a
child rel that is not a baserel. Instead, add a new RelOptInfo
member top_parent_relids to track the same kind of information in a
more generic manner.
Ashutosh Bapat, slightly tweaked by me. Review and testing of the
patch set from which this was taken by Rajkumar Raghuwanshi and Rafia
Sabih.
Discussion: http://postgr.es/m/CA+TgmoagTnF2yqR3PT2rv=om=wJiZ4-A+ATwdnriTGku1CLYxA@mail.gmail.com
gset_data (aka gd) in planner.c is always non-null if and only if
parse->groupingSets is non-null, but Coverity doesn't know that and
complains. Feed it an assertion to see if that keeps it happy.
A QueryEnvironment concept is added, which allows new types of
objects to be passed into queries from parsing on through
execution. At this point, the only thing implemented is a
collection of EphemeralNamedRelation objects -- relations which
can be referenced by name in queries, but do not exist in the
catalogs. The only type of ENR implemented is NamedTuplestore, but
provision is made to add more types fairly easily.
An ENR can carry its own TupleDesc or reference a relation in the
catalogs by relid.
Although these features can be used without SPI, convenience
functions are added to SPI so that ENRs can easily be used by code
run through SPI.
The initial use of all this is going to be transition tables in
AFTER triggers, but that will be added to each PL as a separate
commit.
An incidental effect of this patch is to produce a more informative
error message if an attempt is made to modify the contents of a CTE
from a referencing DML statement. No tests previously covered that
possibility, so one is added.
Kevin Grittner and Thomas Munro
Reviewed by Heikki Linnakangas, David Fetter, and Thomas Munro
with valuable comments and suggestions from many others
Commit 45be99f8cd removed GatherPath's
num_workers field, but this is entirely bogus. Normally, a path's
parallel_workers flag is supposed to indicate the number of workers
that it wants, and should be 0 for a non-partial path. In that
commit, I mistakenly thought that GatherPath could also use that field
to indicate the number of workers that it would try to start, but
that's disastrous, because then it can propagate up to higher nodes in
the plan tree, which will then get incorrect rowcounts because the
parallel_workers flag is involved in computing those values. Repair
by putting the separate field back.
Report by Tomas Vondra. Patch by me, reviewed by Amit Kapila.
Discussion: http://postgr.es/m/f91b4a44-f739-04bd-c4b6-f135bd643669@2ndquadrant.com
copyObject() is declared to return void *, which allows easily assigning
the result independent of the input, but it loses all type checking.
If the compiler supports typeof or something similar, cast the result to
the input type. This creates a greater amount of type safety. In some
cases, where the result is assigned to a generic type such as Node * or
Expr *, new casts are now necessary, but in general casts are now
unnecessary in the normal case and indicate that something unusual is
happening.
Reviewed-by: Mark Dilger <hornschnorter@gmail.com>
This extends the Aggregate node with two new features: HashAggregate
can now run multiple hashtables concurrently, and a new strategy
MixedAggregate populates hashtables while doing sorted grouping.
The planner will now attempt to save as many sorts as possible when
planning grouping sets queries, while not exceeding work_mem for the
estimated combined sizes of all hashtables used. No SQL-level changes
are required. There should be no user-visible impact other than the
new EXPLAIN output and possible changes to result ordering when ORDER
BY was not used (which affected a few regression tests). The
enable_hashagg option is respected.
Author: Andrew Gierth
Reviewers: Mark Dilger, Andres Freund
Discussion: https://postgr.es/m/87vatszyhj.fsf@news-spur.riddles.org.uk
This replaces the old, recursive tree-walk based evaluation, with
non-recursive, opcode dispatch based, expression evaluation.
Projection is now implemented as part of expression evaluation.
This both leads to significant performance improvements, and makes
future just-in-time compilation of expressions easier.
The speed gains primarily come from:
- non-recursive implementation reduces stack usage / overhead
- simple sub-expressions are implemented with a single jump, without
function calls
- sharing some state between different sub-expressions
- reduced amount of indirect/hard to predict memory accesses by laying
out operation metadata sequentially; including the avoidance of
nearly all of the previously used linked lists
- more code has been moved to expression initialization, avoiding
constant re-checks at evaluation time
Future just-in-time compilation (JIT) has become easier, as
demonstrated by released patches intended to be merged in a later
release, for primarily two reasons: Firstly, due to a stricter split
between expression initialization and evaluation, less code has to be
handled by the JIT. Secondly, due to the non-recursive nature of the
generated "instructions", less performance-critical code-paths can
easily be shared between interpreted and compiled evaluation.
The new framework allows for significant future optimizations. E.g.:
- basic infrastructure for to later reduce the per executor-startup
overhead of expression evaluation, by caching state in prepared
statements. That'd be helpful in OLTPish scenarios where
initialization overhead is measurable.
- optimizing the generated "code". A number of proposals for potential
work has already been made.
- optimizing the interpreter. Similarly a number of proposals have
been made here too.
The move of logic into the expression initialization step leads to some
backward-incompatible changes:
- Function permission checks are now done during expression
initialization, whereas previously they were done during
execution. In edge cases this can lead to errors being raised that
previously wouldn't have been, e.g. a NULL array being coerced to a
different array type previously didn't perform checks.
- The set of domain constraints to be checked, is now evaluated once
during expression initialization, previously it was re-built
every time a domain check was evaluated. For normal queries this
doesn't change much, but e.g. for plpgsql functions, which caches
ExprStates, the old set could stick around longer. The behavior
around might still change.
Author: Andres Freund, with significant changes by Tom Lane,
changes by Heikki Linnakangas
Reviewed-By: Tom Lane, Heikki Linnakangas
Discussion: https://postgr.es/m/20161206034955.bh33paeralxbtluv@alap3.anarazel.de
Add support for explicitly declared statistic objects (CREATE
STATISTICS), allowing collection of statistics on more complex
combinations that individual table columns. Companion commands DROP
STATISTICS and ALTER STATISTICS ... OWNER TO / SET SCHEMA / RENAME are
added too. All this DDL has been designed so that more statistic types
can be added later on, such as multivariate most-common-values and
multivariate histograms between columns of a single table, leaving room
for permitting columns on multiple tables, too, as well as expressions.
This commit only adds support for collection of n-distinct coefficient
on user-specified sets of columns in a single table. This is useful to
estimate number of distinct groups in GROUP BY and DISTINCT clauses;
estimation errors there can cause over-allocation of memory in hashed
aggregates, for instance, so it's a worthwhile problem to solve. A new
special pseudo-type pg_ndistinct is used.
(num-distinct estimation was deemed sufficiently useful by itself that
this is worthwhile even if no further statistic types are added
immediately; so much so that another version of essentially the same
functionality was submitted by Kyotaro Horiguchi:
https://postgr.es/m/20150828.173334.114731693.horiguchi.kyotaro@lab.ntt.co.jp
though this commit does not use that code.)
Author: Tomas Vondra. Some code rework by Álvaro.
Reviewed-by: Dean Rasheed, David Rowley, Kyotaro Horiguchi, Jeff Janes,
Ideriha Takeshi
Discussion: https://postgr.es/m/543AFA15.4080608@fuzzy.czhttps://postgr.es/m/20170320190220.ixlaueanxegqd5gr@alvherre.pgsql
Partitioned tables do not contain any data; only their unpartitioned
descendents need to be scanned. However, the partitioned tables still
need to be locked, even though they're not scanned. To make that
work, Append and MergeAppend relations now need to carry a list of
(unscanned) partitioned relations that must be locked, and InitPlan
must lock all partitioned result relations.
Aside from the obvious advantage of avoiding some work at execution
time, this has two other advantages. First, it may improve the
planner's decision-making in some cases since the empty relation
might throw things off. Second, it paves the way to getting rid of
the storage for partitioned tables altogether.
Amit Langote, reviewed by me.
Discussion: http://postgr.es/m/6837c359-45c4-8044-34d1-736756335a15@lab.ntt.co.jp
Partitionwise join proposes add a concept of child join relations,
which will have the same relationship with join relations as "other
member" relations do with base relations. These relations will need
some but not all of the handling that we currently have for join
relations, and some but not all of the handling that we currently have
for appendrels, since they are a mix of the two. Refactor a little
bit so that the necessary bits of logic are exposed as separate
functions.
Ashutosh Bapat, reviewed and tested by Rajkumar Raghuwanshi and
by me.
Discussion: http://postgr.es/m/CAFjFpRfqotRR6cM3sooBHMHEVdkFfAZ6PyYg4GRZsoMuW08HjQ@mail.gmail.com
Commit 51ee6f3160 accidentally changed
the behavior around inheritance hierarchies; before, we always
considered parallel paths even for very small inheritance children,
because otherwise an inheritance hierarchy with even one small child
wouldn't be eligible for parallelism. That exception was inadverently
removed; put it back.
In passing, also adjust the degree-of-parallelism comptuation for
index-only scans not to consider the number of heap pages fetched.
Otherwise, we'll avoid parallel index-only scans on tables that are
mostly all-visible, which isn't especially logical.
Robert Haas and Amit Kapila, per a report from Ashutosh Sharma.
Discussion: http://postgr.es/m/CAE9k0PmgSoOHRd60SHu09aRVTHRSs8s6pmyhJKWHxWw9C_x+XA@mail.gmail.com
This logic was adapated from create_merge_append_plan, but the two
cases aren't really analogous, because create_merge_append_plan is not
projection-capable and must therefore have a tlist identical to that
of the underlying paths. Overwriting the tlist of Gather Merge with
whatever the underlying plan happens to produce is no good at all.
Patch by me, reviewed by Rushabh Lathia, who also reported the issue
and made an initial attempt at a fix.
Discussion: http://postgr.es/m/CA+Tgmob_-oHEOBfT9S25bjqokdqv8e8xEmh9zOY+3MPr_LmuhA@mail.gmail.com
Fallout from fcec6caafa2: mark a variable in
set_tablefunc_size_estimates as used for asserts only.
Also, the planner_rte_fetch() call is pointless with assertions
disabled, so enclose it in a USE_ASSERT_CHECKING #ifdef; fix the same
problem in set_subquery_size_estimates().
First problem noted by David Rowley, whose compiler is noisier than mine
in this regard.
Like Gather, we spawn multiple workers and run the same plan in each
one; however, Gather Merge is used when each worker produces the same
output ordering and we want to preserve that output ordering while
merging together the streams of tuples from various workers. (In a
way, Gather Merge is like a hybrid of Gather and MergeAppend.)
This works out to a win if it saves us from having to perform an
expensive Sort. In cases where only a small amount of data would need
to be sorted, it may actually be faster to use a regular Gather node
and then sort the results afterward, because Gather Merge sometimes
needs to wait synchronously for tuples whereas a pure Gather generally
doesn't. But if this avoids an expensive sort then it's a win.
Rushabh Lathia, reviewed and tested by Amit Kapila, Thomas Munro,
and Neha Sharma, and reviewed and revised by me.
Discussion: http://postgr.es/m/CAGPqQf09oPX-cQRpBKS0Gq49Z+m6KBxgxd_p9gX8CKk_d75HoQ@mail.gmail.com
The index is scanned by a single process, but then all cooperating
processes can iterate jointly over the resulting set of heap blocks.
In the future, we might also want to support using a parallel bitmap
index scan to set up for a parallel bitmap heap scan, but that's a
job for another day.
Dilip Kumar, with some corrections and cosmetic changes by me. The
larger patch set of which this is a part has been reviewed and tested
by (at least) Andres Freund, Amit Khandekar, Tushar Ahuja, Rafia
Sabih, Haribabu Kommi, Thomas Munro, and me.
Discussion: http://postgr.es/m/CAFiTN-uc4=0WxRGfCzs-xfkMYcSEWUC-Fon6thkJGjkh9i=13A@mail.gmail.com
XMLTABLE is defined by the SQL/XML standard as a feature that allows
turning XML-formatted data into relational form, so that it can be used
as a <table primary> in the FROM clause of a query.
This new construct provides significant simplicity and performance
benefit for XML data processing; what in a client-side custom
implementation was reported to take 20 minutes can be executed in 400ms
using XMLTABLE. (The same functionality was said to take 10 seconds
using nested PostgreSQL XPath function calls, and 5 seconds using
XMLReader under PL/Python).
The implemented syntax deviates slightly from what the standard
requires. First, the standard indicates that the PASSING clause is
optional and that multiple XML input documents may be given to it; we
make it mandatory and accept a single document only. Second, we don't
currently support a default namespace to be specified.
This implementation relies on a new executor node based on a hardcoded
method table. (Because the grammar is fixed, there is no extensibility
in the current approach; further constructs can be implemented on top of
this such as JSON_TABLE, but they require changes to core code.)
Author: Pavel Stehule, Álvaro Herrera
Extensively reviewed by: Craig Ringer
Discussion: https://postgr.es/m/CAFj8pRAgfzMD-LoSmnMGybD0WsEznLHWap8DO79+-GTRAPR4qA@mail.gmail.com
Commit 45be99f8cd took the position
that performing a merge join in parallel was not likely to work out
well, but this conclusion was greeted with skepticism even at the
time. Whether it was true then or not, it's clearly not true any
more now that we have parallel index scan.
Dilip Kumar, reviewed by Amit Kapila and by me.
Discussion: http://postgr.es/m/CAFiTN-v3=cM6nyFwFGp0fmvY4=kk79Hq9Fgu0u8CSJ-EEq1Tiw@mail.gmail.com
Extract the logic used by hash_inner_and_outer into a separate
function, get_cheapest_parallel_safe_total_inner, so that it can
also be used to plan parallel merge joins.
Also, add a require_parallel_safe argument to the existing function
get_cheapest_path_for_pathkeys, because parallel merge join needs
to find the cheapest path for a given set of pathkeys that is
parallel-safe, not just the cheapest one overall.
Patch by me, reviewed by Dilip Kumar.
Discussion: http://postgr.es/m/CA+TgmoYOv+dFK0MWW6366dFj_xTnohQfoBDrHyB7d1oZhrgPjA@mail.gmail.com
When the very cheapest path is not parallel-safe, we want to instead use
the cheapest unparameterized path that is. The old code searched
innerrel->cheapest_parameterized_paths, but that isn't right, because
the path we want may not be in that list. Search innerrel->pathlist
instead.
Spotted by Dilip Kumar.
Discussion: http://postgr.es/m/CAFiTN-szCEcZrQm0i_w4xqSaRUTOUFstNu32Zn4rxxDcoa8gnA@mail.gmail.com
Commit 5262f7a4fc added similar support
for parallel index scans; this extends that work to index-only scans.
As with parallel index scans, this requires support from the index AM,
so currently parallel index-only scans will only be possible for btree
indexes.
Rafia Sabih, reviewed and tested by Rahila Syed, Tushar Ahuja,
and Amit Kapila
Discussion: http://postgr.es/m/CAOGQiiPEAs4C=TBp0XShxBvnWXuzGL2u++Hm1=qnCpd6_Mf8Fw@mail.gmail.com
In combination with 569174f1be, which
taught the btree AM how to perform parallel index scans, this allows
parallel index scan plans on btree indexes. This infrastructure
should be general enough to support parallel index scans for other
index AMs as well, if someone updates them to support parallel
scans.
Amit Kapila, reviewed and tested by Anastasia Lubennikova, Tushar
Ahuja, and Haribabu Kommi, and me.
When min_parallel_relation_size was added, the only supported type
of parallel scan was a parallel sequential scan, but there are
pending patches for parallel index scan, parallel index-only scan,
and parallel bitmap heap scan. Those patches introduce two new
types of complications: first, what's relevant is not really the
total size of the relation but the portion of it that we will scan;
and second, index pages and heap pages shouldn't necessarily be
treated in exactly the same way. Typically, the number of index
pages will be quite small, but that doesn't necessarily mean that
a parallel index scan can't pay off.
Therefore, we introduce min_parallel_table_scan_size, which works
out a degree of parallelism for scans based on the number of table
pages that will be scanned (and which is therefore equivalent to
min_parallel_relation_size for parallel sequential scans) and also
min_parallel_index_scan_size which can be used to work out a degree
of parallelism based on the number of index pages that will be
scanned.
Amit Kapila and Robert Haas
Discussion: http://postgr.es/m/CAA4eK1KowGSYYVpd2qPpaPPA5R90r++QwDFbrRECTE9H_HvpOg@mail.gmail.com
Discussion: http://postgr.es/m/CAA4eK1+TnM4pXQbvn7OXqam+k_HZqb0ROZUMxOiL6DWJYCyYow@mail.gmail.com
This doesn't do anything to make Param nodes anything other than
parallel-restricted, so this only helps with uncorrelated subplans,
and it's not necessarily very cheap because each worker will run the
subplan separately (just as a Hash Join will build a separate copy of
the hash table in each participating process), but it's a first step
toward supporting cases that are more likely to help in practice, and
is occasionally useful on its own.
Amit Kapila, reviewed and tested by Rafia Sabih, Dilip Kumar, and
me.
Discussion: http://postgr.es/m/CAA4eK1+e8Z45D2n+rnDMDYsVEb5iW7jqaCH_tvPMYau=1Rru9w@mail.gmail.com
I noticed while hacking on join UNION transforms that planner.c's
function get_base_rel_indexes() just duplicates the functionality of
get_relids_in_jointree(). It doesn't even have the excuse of being
older code :-(. Drop it and use the latter function instead.
If we forcibly place a Material node atop a finished subplan, we need
to move any initPlans attached to the subplan up to the Material node,
in order to keep SS_finalize_plan() happy. I'd figured this out in
commit 7b67a0a49 for the case of materializing a cursor plan, but out of
an abundance of caution, I put the initPlan movement hack at the call
site for that case, rather than inside materialize_finished_plan().
That was the wrong thing, because it turns out to also be necessary for
the only other caller of materialize_finished_plan(), ie subselect.c.
We lacked any test cases that exposed the mistake, but bug#14524 from
Wei Congrui shows that it's possible to get an initPlan reference into
the top tlist in that case too, and then SS_finalize_plan() complains.
Hence, move the hack into materialize_finished_plan().
In HEAD, also relocate some recently-added tests in subselect.sql, which
I'd unthinkingly dropped into the middle of a sequence of related tests.
Report: https://postgr.es/m/20170202060020.1400.89021@wrigleys.postgresql.org
Given a targetlist like "srf(x), f(srf(x))", split_pathtarget_at_srfs()
decided that it needed two levels of ProjectSet nodes, failing to notice
that the two SRF calls are textually equal(). Because of that, setrefs.c
would convert the upper ProjectSet's tlist to "Var1, f(Var1)" (where Var1
represents a reference to the srf(x) output of the lower ProjectSet).
This triggered an assertion in nodeProjectSet.c complaining that it found
no SRFs to evaluate, as reported by Erik Rijkers.
What we want in such a case is to evaluate srf(x) only once and use a plain
Result node to compute "Var1, f(Var1)"; that gives results similar to what
previous versions produced, whereas allowing srf(x) to be evaluated again
in an upper ProjectSet would square the number of rows emitted.
Furthermore, even if the SRF calls aren't textually identical, we want them
to be evaluated in lockstep, because that's what happened in the old
implementation. But split_pathtarget_at_srfs() got this completely wrong,
using two levels of ProjectSet for a case like "srf(x), f(srf(y))".
Hence, rewrite split_pathtarget_at_srfs() from the ground up so that it
groups SRFs according to the depth of nesting of SRFs in their arguments.
This is pretty much how we envisioned that working originally, but I blew
it when it came to implementation.
In passing, optimize the case of target == input_target, which I noticed
is not only possible but quite common.
Discussion: https://postgr.es/m/dcbd2853c05d22088766553d60dc78c6@xs4all.nl
Currently, we only need this logic in order to cost a Bitmap Heap
Scan. But a pending patch for Parallel Bitmap Heap Scan also uses
it to help figure out how many workers to use for the scan, which
has to be determined prior to costing. So, move the logic to
a separate function to make that easier.
Dilip Kumar. The patch series of which this is a part has been
reviewed by Andres Freund, Amit Khendekar, Tushar Ahuja, Rafia
Sabih, Haribabu Kommi, and me; it is not clear from the email
discussion which of those people have looked specifically at this
part.
Discussion: http://postgr.es/m/CAFiTN-v3QYNJEZnnmKCeATuLbN-h9tMVfeEF0+BrouYDqjXgwg@mail.gmail.com
Joining three tables only takes two join nodes. I think when I (tgl)
wrote this, I was envisioning possible additional joins; but since the
example doesn't show any fourth table, it's just confusing to write
a third join node.
Etsuro Fujita
Discussion: https://postgr.es/m/e6cfbaa3-af02-1abc-c25e-8fa5c6bc4e21@lab.ntt.co.jp
I'd somehow talked myself into believing that set_append_rel_size
doesn't need to worry about getting back an AND clause when it applies
eval_const_expressions to the result of adjust_appendrel_attrs (that is,
transposing the appendrel parent's restriction clauses for one child).
But that is nonsense, and Andreas Seltenreich's fuzz tester soon
turned up a counterexample. Put back the make_ands_implicit step
that was there before, and add a regression test covering the case.
Report: https://postgr.es/m/878tq6vja6.fsf@ansel.ydns.eu
Since 69f4b9c plain expression evaluation (and thus normal projection)
can't return sets of tuples anymore. Thus remove code dealing with
that possibility.
This will require adjustments in external code using
ExecEvalExpr()/ExecProject() - that should neither be hard nor very
common.
Author: Andres Freund and Tom Lane
Discussion: https://postgr.es/m/20160822214023.aaxz5l4igypowyri@alap3.anarazel.de
Evaluation of set returning functions (SRFs_ in the targetlist (like SELECT
generate_series(1,5)) so far was done in the expression evaluation (i.e.
ExecEvalExpr()) and projection (i.e. ExecProject/ExecTargetList) code.
This meant that most executor nodes performing projection, and most
expression evaluation functions, had to deal with the possibility that an
evaluated expression could return a set of return values.
That's bad because it leads to repeated code in a lot of places. It also,
and that's my (Andres's) motivation, made it a lot harder to implement a
more efficient way of doing expression evaluation.
To fix this, introduce a new executor node (ProjectSet) that can evaluate
targetlists containing one or more SRFs. To avoid the complexity of the old
way of handling nested expressions returning sets (e.g. having to pass up
ExprDoneCond, and dealing with arguments to functions returning sets etc.),
those SRFs can only be at the top level of the node's targetlist. The
planner makes sure (via split_pathtarget_at_srfs()) that SRF evaluation is
only necessary in ProjectSet nodes and that SRFs are only present at the
top level of the node's targetlist. If there are nested SRFs the planner
creates multiple stacked ProjectSet nodes. The ProjectSet nodes always get
input from an underlying node.
We also discussed and prototyped evaluating targetlist SRFs using ROWS
FROM(), but that turned out to be more complicated than we'd hoped.
While moving SRF evaluation to ProjectSet would allow to retain the old
"least common multiple" behavior when multiple SRFs are present in one
targetlist (i.e. continue returning rows until all SRFs are at the end of
their input at the same time), we decided to instead only return rows till
all SRFs are exhausted, returning NULL for already exhausted ones. We
deemed the previous behavior to be too confusing, unexpected and actually
not particularly useful.
As a side effect, the previously prohibited case of multiple set returning
arguments to a function, is now allowed. Not because it's particularly
desirable, but because it ends up working and there seems to be no argument
for adding code to prohibit it.
Currently the behavior for COALESCE and CASE containing SRFs has changed,
returning multiple rows from the expression, even when the SRF containing
"arm" of the expression is not evaluated. That's because the SRFs are
evaluated in a separate ProjectSet node. As that's quite confusing, we're
likely to instead prohibit SRFs in those places. But that's still being
discussed, and the code would reside in places not touched here, so that's
a task for later.
There's a lot of, now superfluous, code dealing with set return expressions
around. But as the changes to get rid of those are verbose largely boring,
it seems better for readability to keep the cleanup as a separate commit.
Author: Tom Lane and Andres Freund
Discussion: https://postgr.es/m/20160822214023.aaxz5l4igypowyri@alap3.anarazel.de
In an RLS query, we must ensure that security filter quals are evaluated
before ordinary query quals, in case the latter contain "leaky" functions
that could expose the contents of sensitive rows. The original
implementation of RLS planning ensured this by pushing the scan of a
secured table into a sub-query that it marked as a security-barrier view.
Unfortunately this results in very inefficient plans in many cases, because
the sub-query cannot be flattened and gets planned independently of the
rest of the query.
To fix, drop the use of sub-queries to enforce RLS qual order, and instead
mark each qual (RestrictInfo) with a security_level field establishing its
priority for evaluation. Quals must be evaluated in security_level order,
except that "leakproof" quals can be allowed to go ahead of quals of lower
security_level, if it's helpful to do so. This has to be enforced within
the ordering of any one list of quals to be evaluated at a table scan node,
and we also have to ensure that quals are not chosen for early evaluation
(i.e., use as an index qual or TID scan qual) if they're not allowed to go
ahead of other quals at the scan node.
This is sufficient to fix the problem for RLS quals, since we only support
RLS policies on simple tables and thus RLS quals will always exist at the
table scan level only. Eventually these qual ordering rules should be
enforced for join quals as well, which would permit improving planning for
explicit security-barrier views; but that's a task for another patch.
Note that FDWs would need to be aware of these rules --- and not, for
example, send an insecure qual for remote execution --- but since we do
not yet allow RLS policies on foreign tables, the case doesn't arise.
This will need to be addressed before we can allow such policies.
Patch by me, reviewed by Stephen Frost and Dean Rasheed.
Discussion: https://postgr.es/m/8185.1477432701@sss.pgh.pa.us
Normally, if we have a WHERE clause like "indexcol = constant",
the planner will figure out that that index column can be ignored
when determining whether the index has a desired sort ordering.
But this failed to work for boolean index columns, because a
condition like "boolcol = true" is canonicalized to just "boolcol"
which does not give rise to an EquivalenceClass. Add a check to
allow the same type of deduction to be made in this case too.
Per a complaint from Dima Pavlov. Arguably this is a bug, but given the
limited impact and the small number of complaints so far, I won't risk
destabilizing plans in stable branches by back-patching.
Patch by me, reviewed by Michael Paquier
Discussion: https://postgr.es/m/1788.1481605684@sss.pgh.pa.us
This patch makes several changes that improve the consistency of
representation of lists of statements. It's always been the case
that the output of parse analysis is a list of Query nodes, whatever
the types of the individual statements in the list. This patch brings
similar consistency to the outputs of raw parsing and planning steps:
* The output of raw parsing is now always a list of RawStmt nodes;
the statement-type-dependent nodes are one level down from that.
* The output of pg_plan_queries() is now always a list of PlannedStmt
nodes, even for utility statements. In the case of a utility statement,
"planning" just consists of wrapping a CMD_UTILITY PlannedStmt around
the utility node. This list representation is now used in Portal and
CachedPlan plan lists, replacing the former convention of intermixing
PlannedStmts with bare utility-statement nodes.
Now, every list of statements has a consistent head-node type depending
on how far along it is in processing. This allows changing many places
that formerly used generic "Node *" pointers to use a more specific
pointer type, thus reducing the number of IsA() tests and casts needed,
as well as improving code clarity.
Also, the post-parse-analysis representation of DECLARE CURSOR is changed
so that it looks more like EXPLAIN, PREPARE, etc. That is, the contained
SELECT remains a child of the DeclareCursorStmt rather than getting flipped
around to be the other way. It's now true for both Query and PlannedStmt
that utilityStmt is non-null if and only if commandType is CMD_UTILITY.
That allows simplifying a lot of places that were testing both fields.
(I think some of those were just defensive programming, but in many places,
it was actually necessary to avoid confusing DECLARE CURSOR with SELECT.)
Because PlannedStmt carries a canSetTag field, we're also able to get rid
of some ad-hoc rules about how to reconstruct canSetTag for a bare utility
statement; specifically, the assumption that a utility is canSetTag if and
only if it's the only one in its list. While I see no near-term need for
relaxing that restriction, it's nice to get rid of the ad-hocery.
The API of ProcessUtility() is changed so that what it's passed is the
wrapper PlannedStmt not just the bare utility statement. This will affect
all users of ProcessUtility_hook, but the changes are pretty trivial; see
the affected contrib modules for examples of the minimum change needed.
(Most compilers should give pointer-type-mismatch warnings for uncorrected
code.)
There's also a change in the API of ExplainOneQuery_hook, to pass through
cursorOptions instead of expecting hook functions to know what to pick.
This is needed because of the DECLARE CURSOR changes, but really should
have been done in 9.6; it's unlikely that any extant hook functions
know about using CURSOR_OPT_PARALLEL_OK.
Finally, teach gram.y to save statement boundary locations in RawStmt
nodes, and pass those through to Query and PlannedStmt nodes. This allows
more intelligent handling of cases where a source query string contains
multiple statements. This patch doesn't actually do anything with the
information, but a follow-on patch will. (Passing this information through
cleanly is the true motivation for these changes; while I think this is all
good cleanup, it's unlikely we'd have bothered without this end goal.)
catversion bump because addition of location fields to struct Query
affects stored rules.
This patch is by me, but it owes a good deal to Fabien Coelho who did
a lot of preliminary work on the problem, and also reviewed the patch.
Discussion: https://postgr.es/m/alpine.DEB.2.20.1612200926310.29821@lancre
For a partial path, the cardinality estimate needs to reflect the
number of rows we think each worker will see, rather than the total
number of rows; otherwise, costing will go wrong. The previous coding
got this completely wrong for parallel joins.
Unfortunately, this change may destabilize plans for users of 9.6 who
have enabled parallel query, but since 9.6 is still fairly new I'm
hoping expectations won't be too settled yet. Also, this is really a
brown-paper-bag bug, so leaving it unfixed for the entire lifetime of
9.6 seems unwise.
Related reports (whose import I initially failed to recognize) by
Tomas Vondra and Tom Lane.
Discussion: http://postgr.es/m/CA+TgmoaDxZ5z5Kw_oCQoymNxNoVaTCXzPaODcOuao=CzK8dMZw@mail.gmail.com
RelationGetPartitionQual() and generate_partition_qual() are always
called with recurse = true, so we don't need an argument for that.
Extracted by me from a larger patch by Amit Langote.
This shouldn't change the set of paths that get generated in any
way, but it is preparatory work for further changes to allow a
partial path to be merge-joined witih a non-partial path to produce
a partial join path.
Dilip Kumar, with cosmetic adjustments by me.
This case wasn't thought through sufficiently in commit 100340e2d.
It's true that the FK proves that every outer row has a match in the
inner table, but we forgot that some of the inner rows might be filtered
away by WHERE conditions located within the semijoin's RHS.
If the RHS is just one table, we can reasonably take the semijoin
selectivity as equal to the fraction of the referenced table's rows
that are expected to survive its restriction clauses.
If the RHS is a join, it's not clear how much of the referenced table
might get through the join, so fall back to the same rule we were
already using for other outer-join cases: use the minimum of the
regular per-clause selectivity estimates. This gives the same result
as if we hadn't considered the FK at all when there's a single FK
column, but it should still help for multi-column FKs, which is the
case that 100340e2d is really meant to help with.
Back-patch to 9.6 where the previous commit came in.
Discussion: https://postgr.es/m/16149.1481835103@sss.pgh.pa.us
The existing tests in preprocess_minmax_aggregates() usually prevent it
from trying to do anything with queries containing CTEs, but there's an
exception: a CTE could be present as a member of an appendrel, if we
flattened a UNION ALL that contains CTE references. If it did try to
generate an optimized path for a query using a CTE, it failed with
"could not find plan for CTE", as reported by Torsten Förtsch.
The proximate cause is an unwise decision in commit 3fc6e2d7f to clear
subroot->cte_plan_ids in build_minmax_path(). That left the subroot's
cte_plan_ids list out of step with its parse->cteList.
Removing the "subroot->cte_plan_ids = NIL;" assignment is enough to let
the case work again, but really it's pretty silly to be expending any
cycles at all in this module when there are CTEs: we always treat their
outputs as unordered so there's no way for the optimization to win.
Hence, also add an early-exit test so we don't waste time like that.
Back-patch to 9.6 where the misbehavior was introduced.
Report: https://postgr.es/m/CAKkG4_=gjY5QiHtqSZyWMwDuTd_CftKoTaCqxjJ7uUz1-Gw=qw@mail.gmail.com
expandRTE() and get_rte_attribute_type() reported the exprType() and
exprTypmod() values of the expressions in the first row of the VALUES as
being the column type/typmod returned by the VALUES RTE. That's fine for
the data type, since we coerce all expressions in a column to have the same
common type. But we don't coerce them to have a common typmod, so it was
possible for rows after the first one to return values that violate the
claimed column typmod. This leads to the incorrect result seen in bug
#14448 from Hassan Mahmood, as well as some other corner-case misbehaviors.
The desired behavior is the same as we use in other type-unification
cases: report the common typmod if there is one, but otherwise return -1
indicating no particular constraint. It's cheap for transformValuesClause
to determine the common typmod while transforming a multi-row VALUES, but
it'd be less cheap for expandRTE() and get_rte_attribute_type() to
re-determine that info every time they're asked --- possibly a lot less
cheap, if the VALUES has many rows. Therefore, the best fix is to record
the common typmods explicitly in a list in the VALUES RTE, as we were
already doing for column collations. This looks quite a bit like what
we're doing for CTE RTEs, so we can save a little bit of space and code by
unifying the representation for those two RTE types. They both now share
coltypes/coltypmods/colcollations fields. (At some point it might seem
desirable to populate those fields for all RTE types; but right now it
looks like constructing them for other RTE types would add more code and
cycles than it would save.)
The RTE change requires a catversion bump, so this fix is only usable
in HEAD. If we fix this at all in the back branches, the patch will
need to look quite different.
Report: https://postgr.es/m/20161205143037.4377.60754@wrigleys.postgresql.org
Discussion: https://postgr.es/m/27429.1480968538@sss.pgh.pa.us
Table partitioning is like table inheritance and reuses much of the
existing infrastructure, but there are some important differences.
The parent is called a partitioned table and is always empty; it may
not have indexes or non-inherited constraints, since those make no
sense for a relation with no data of its own. The children are called
partitions and contain all of the actual data. Each partition has an
implicit partitioning constraint. Multiple inheritance is not
allowed, and partitioning and inheritance can't be mixed. Partitions
can't have extra columns and may not allow nulls unless the parent
does. Tuples inserted into the parent are automatically routed to the
correct partition, so tuple-routing ON INSERT triggers are not needed.
Tuple routing isn't yet supported for partitions which are foreign
tables, and it doesn't handle updates that cross partition boundaries.
Currently, tables can be range-partitioned or list-partitioned. List
partitioning is limited to a single column, but range partitioning can
involve multiple columns. A partitioning "column" can be an
expression.
Because table partitioning is less general than table inheritance, it
is hoped that it will be easier to reason about properties of
partitions, and therefore that this will serve as a better foundation
for a variety of possible optimizations, including query planner
optimizations. The tuple routing based which this patch does based on
the implicit partitioning constraints is an example of this, but it
seems likely that many other useful optimizations are also possible.
Amit Langote, reviewed and tested by Robert Haas, Ashutosh Bapat,
Amit Kapila, Rajkumar Raghuwanshi, Corey Huinker, Jaime Casanova,
Rushabh Lathia, Erik Rijkers, among others. Minor revisions by me.
consider_parallel_nestloop passed the wrong jointype down to its
subroutines for JOIN_UNIQUE_INNER cases (it should pass JOIN_INNER), and it
thought that it could pass paths other than innerrel->cheapest_total_path
to create_unique_path, which create_unique_path is not on board with.
These bugs would lead to assertion failures or other errors, suggesting
that this code path hasn't been tested much.
hash_inner_and_outer's code for parallel join effectively treated both
JOIN_UNIQUE_OUTER and JOIN_UNIQUE_INNER the same as JOIN_INNER (for
different reasons :-(), leading to incorrect plans that treated a semijoin
as if it were a plain join.
Michael Day submitted a test case demonstrating that hash_inner_and_outer
failed for JOIN_UNIQUE_OUTER, and I found the other cases through code
review.
Report: https://postgr.es/m/D0E8A029-D1AC-42E8-979A-5DE4A77E4413@rcmail.com
Oversight in my commit 0bb51aa96. Noted while poking at a recent
bug report --- HEAD's estimates for a query using CURRENT_DATE
were unexpectedly much worse than 9.6's.
Andreas Seltenreich found another case where we were being too optimistic
about allowing a plan to be considered parallelizable despite it containing
initPlans. It seems like the real issue here is that if we know we are
going to tack initPlans onto the topmost Plan node for a subquery, we
had better mark that subquery's result Paths as not-parallel-safe. That
fixes this problem and allows reversion of a kluge (added in commit
7b67a0a49 and extended in f24cf960d) to not trust the parallel_safe flag
at top level.
Discussion: <874m2w4k5d.fsf@ex.ansel.ydns.eu>
Fix thinko in commit da1c91631: even if the original query was free of
parallel hazards, we might introduce such a hazard by adding PARAM_EXEC
Param nodes. Adjust is_parallel_safe() so that it will scan the given
expression whenever any such nodes have been created. Per report from
Andreas Seltenreich.
Discussion: <878tse6yvf.fsf@credativ.de>
We mustn't force parallel mode if the query has any subplans, since
ExecSerializePlan doesn't transmit them to workers. Testing
top_plan->initPlan is inadequate because (1) there might be initPlans
attached to lower plan nodes, and (2) non-initPlan subplans don't
work either. There's certainly room for improvement in those
restrictions, but for the moment that's what we've got.
Amit Kapila, per report from Andreas Seltenreich
Discussion: <8737im6pmh.fsf@credativ.de>
create_merge_append_path should use the path rowcount it just computed,
not rel->tuples, for costing purposes. Those numbers should always be
the same at present, but if we ever support parameterized MergeAppend
paths (a case this function is otherwise prepared for), the former would
be right and the latter wrong.
No need for back-patch since the problem is only latent.
Ashutosh Bapat
Discussion: <CAFjFpRek+cLCnTo24youuGtsq4zRphEB8EUUPjDxZjnL4n4HYQ@mail.gmail.com>
Be sure to pull up the subquery's hasRowSecurity flag when flattening a
subquery in pull_up_simple_subquery(). This isn't a bug today because
we don't look at the hasRowSecurity flag during planning, but it could
easily be a bug tomorrow.
Likewise, make rewriteRuleAction() pull up the hasRowSecurity flag when
absorbing RTEs from a rule action. This isn't a bug either, for the
opposite reason: the flag should never be set yet. But again, it seems
like good future proofing.
Add a comment explaining why rewriteTargetView() should *not* set
hasRowSecurity when adding stuff to securityQuals.
Improve some nearby comments about securityQuals processing, and document
that field more completely in parsenodes.h.
Patch by me, analysis by Dean Rasheed.
Discussion: <CAEZATCXZ8tb2DV6f=bkhsMV6u_gRcZ0CZBw2J-qU84RxSukZog@mail.gmail.com>
The plan generated for sorted partial aggregation with "GROUP BY constant"
included a Sort node with no sort keys, which the executor does not like.
Per report from Steve Randall. I'd add a regression test case if I could
think of a compact one, but it doesn't seem worth expending lots of cycles
on.
Report: <CABVd52UAdGXpg_rCk46egpNKYdXOzCjuJ1zG26E2xBe_8bj+Fg@mail.gmail.com>
Commit 7a2fe9bd0 improved merge append so that replacement of a tuple
takes log(N) operations, not twice log(N). Since cost_merge_append knew
about that explicitly, we should adjust it. This probably makes little
difference in practice, but the obsolete comment is confusing.
Ideally this would have been put in in 9.3 with the underlying behavior
change; but I'm not going to back-patch it, since there's some small chance
of changing a plan choice that somebody's optimized for.
Thomas Munro
Discussion: <CAEepm=0WQBSvuYcMOUj4Ga4NXpu2J=ejZcE=e=eiTjTX-6_gDw@mail.gmail.com>
The foreign-key-aware logic for estimation of join sizes (added in commit
100340e2d) blindly tried to apply the concept to rels that are actually
parents of inheritance trees. This is just plain wrong so far as the
referenced relation is concerned, since the inheritance scan may well
produce lots of rows that are not participating in the constraint. It's
wrong for the referencing relation too, for the same reason; although on
that end we could conceivably detect whether all members of the inheritance
tree have equivalent FK constraints pointing to the same referenced rel,
and then proceed more or less as we do now. But pending somebody writing
code to do that, we must disable this, because it's producing completely
silly estimates when there's an FK linking the heads of inheritance trees.
Per bug #14404 from Clinton Adams. Back-patch to 9.6 where the new
estimation logic came in.
Report: <20161028200412.15987.96482@wrigleys.postgresql.org>
While converting expressions in an upper-level plan node so that they
reference Vars and expressions provided by the input plan node(s),
don't convert plain Const items, even if there happens to be a matching
Const in the input. It's silly to do so because a Var is more expensive to
execute than a Const. Moreover, converting can fool ExecCheckPlanOutput's
check that an insert or update query inserts nulls into dropped columns,
leading to "query provides a value for a dropped column" errors during
INSERT or UPDATE on a table with a dropped column. We could solve this
by making that check more complicated, but I don't see the point; this fix
should save a marginal number of cycles, and it also makes for less messy
EXPLAIN output, as shown by the ensuing regression test result changes.
Per report from Pavel Hanák. I have not incorporated a test case based
on that example, as there doesn't seem to be a simple way of checking
this in isolation without making a bunch of assumptions about other
planner and SQL-function behavior.
Back-patch to 9.6. This setrefs.c behavior exists much further back,
but there is not currently reason to think that it causes problems
before 9.6.
Discussion: <83shraampf.fsf@is-it.eu>
In the previous coding, if an aggregate's transition function returned an
expanded array, nodeAgg.c and nodeWindowAgg.c would always copy it and thus
force it into the flat representation. This led to ping-ponging between
flat and expanded formats, which costs a lot. For an aggregate using
array_append as transition function, I measured about a 15X slowdown
compared to the pre-9.5 code, when working on simple int[] arrays.
Of course, the old code was already O(N^2) in this usage due to copying
flat arrays all the time, but it wasn't quite this inefficient.
To fix, teach nodeAgg.c and nodeWindowAgg.c to allow expanded transition
values without copying, so long as the transition function takes care to
return the transition value already properly parented under the aggcontext.
That puts a bit of extra responsibility on the transition function, but
doing it this way allows us to not need any extra logic in the fast path
of advance_transition_function (ie, with a pass-by-value transition value,
or with a modified-in-place pass-by-reference value). We already know
that that's a hot spot so I'm loath to add any cycles at all there. Also,
while only array_append currently knows how to follow this convention,
this solution allows other transition functions to opt-in without needing
to have a whitelist in the core aggregation code.
(The reason we would need a whitelist is that currently, if you pass a
R/W expanded-object pointer to an arbitrary function, it's allowed to do
anything with it including deleting it; that breaks the core agg code's
assumption that it should free discarded values. Returning a value under
aggcontext is the transition function's signal that it knows it is an
aggregate transition function and will play nice. Possibly the API rules
for expanded objects should be refined, but that would not be a
back-patchable change.)
With this fix, an aggregate using array_append is no longer O(N^2), so it's
much faster than pre-9.5 code rather than much slower. It's still a bit
slower than the bespoke infrastructure for array_agg, but the differential
seems to be only about 10%-20% rather than orders of magnitude.
Discussion: <6315.1477677885@sss.pgh.pa.us>
Now that the upper planner uses paths, and now that we have proper hooks
to inject paths into the upper planning process, it's possible for
foreign data wrappers to arrange to push aggregates to the remote side
instead of fetching all of the rows and aggregating them locally. This
figures to be a massive win for performance, so teach postgres_fdw to
do it.
Jeevan Chalke and Ashutosh Bapat. Reviewed by Ashutosh Bapat with
additional testing by Prabhat Sahu. Various mostly cosmetic changes
by me.
The more efficient hashtable speeds up hash-aggregations with more than
a few hundred groups significantly. Improvements of over 120% have been
measured.
Due to the the different hash table queries that not fully
determined (e.g. GROUP BY without ORDER BY) may change their result
order.
The conversion is largely straight-forward, except that, due to the
static element types of simplehash.h type hashes, the additional data
some users store in elements (e.g. the per-group working data for hash
aggregaters) is now stored in TupleHashEntryData->additional. The
meaning of BuildTupleHashTable's entrysize (renamed to additionalsize)
has been changed to only be about the additionally stored size. That
size is only used for the initial sizing of the hash-table.
Reviewed-By: Tomas Vondra
Discussion: <20160727004333.r3e2k2y6fvk2ntup@alap3.anarazel.de>
Pushing an upper-level restriction clause into an unflattened
subquery-in-FROM is okay when the subquery contains no SRFs in its
targetlist, or when it does but the SRFs are unreferenced by the clause
*and the clause is not volatile*. Otherwise, we're changing the number
of times the clause is evaluated, which is bad for volatile quals, and
possibly changing the result, since a volatile qual might succeed for some
SRF output rows and not others despite not referencing any of the changing
columns. (Indeed, if the clause is something like "random() > 0.5", the
user is probably expecting exactly that behavior.)
We had most of these restrictions down, but not the one about the upper
clause not being volatile. Fix that, and add a regression test to
illustrate the expected behavior.
Although this is definitely a bug, it doesn't seem like back-patch
material, since possibly some users don't realize that the broken
behavior is broken and are relying on what happens now. Also, while
the added test is quite cheap in the wake of commit a4c35ea1c, it would
be much more expensive (or else messier) in older branches.
Per report from Tom van Tilburg.
Discussion: <CAP3PPDiucxYCNev52=YPVkrQAPVF1C5PFWnrQPT7iMzO1fiKFQ@mail.gmail.com>
Teach the parser to reject misplaced set-returning functions during parse
analysis using p_expr_kind, in much the same way as we do for aggregates
and window functions (cf commit eaccfded9). While this isn't complete
(it misses nesting-based restrictions), it's much better than the previous
error reporting for such cases, and it allows elimination of assorted
ad-hoc expression_returns_set() error checks. We could add nesting checks
later if it seems important to catch all cases at parse time.
There is one case the parser will now throw error for although previous
versions allowed it, which is SRFs in the tlist of an UPDATE. That never
behaved sensibly (since it's ill-defined which generated row should be
used to perform the update) and it's hard to see why it should not be
treated as an error. It's a release-note-worthy change though.
Also, add a new Query field hasTargetSRFs reporting whether there are
any SRFs in the targetlist (including GROUP BY/ORDER BY expressions).
The parser can now set that basically for free during parse analysis,
and we can use it in a number of places to avoid expression_returns_set
searches. (There will be more such checks soon.) In some places, this
allows decontorting the logic since it's no longer expensive to check for
SRFs in the tlist --- so I made the checks parallel to the handling of
hasAggs/hasWindowFuncs wherever it seemed appropriate.
catversion bump because adding a Query field changes stored rules.
Andres Freund and Tom Lane
Discussion: <24639.1473782855@sss.pgh.pa.us>
I found that half a dozen (nearly 5%) of our AllocSetContextCreate calls
had typos in the context-sizing parameters. While none of these led to
especially significant problems, they did create minor inefficiencies,
and it's now clear that expecting people to copy-and-paste those calls
accurately is not a great idea. Let's reduce the risk of future errors
by introducing single macros that encapsulate the common use-cases.
Three such macros are enough to cover all but two special-purpose contexts;
those two calls can be left as-is, I think.
While this patch doesn't in itself improve matters for third-party
extensions, it doesn't break anything for them either, and they can
gradually adopt the simplified notation over time.
In passing, change TopMemoryContext to use the default allocation
parameters. Formerly it could only be extended 8K at a time. That was
probably reasonable when this code was written; but nowadays we create
many more contexts than we did then, so that it's not unusual to have a
couple hundred K in TopMemoryContext, even without considering various
dubious code that sticks other things there. There seems no good reason
not to let it use growing blocks like most other contexts.
Back-patch to 9.6, mostly because that's still close enough to HEAD that
it's easy to do so, and keeping the branches in sync can be expected to
avoid some future back-patching pain. The bugs fixed by these changes
don't seem to be significant enough to justify fixing them further back.
Discussion: <21072.1472321324@sss.pgh.pa.us>
ExecReScanAgg's check for whether it could re-use a previously calculated
hashtable neglected the possibility that the Agg node might reference
PARAM_EXEC Params that are not referenced by its input plan node. That's
okay if the Params are in upper tlist or qual expressions; but if one
appears in aggregate input expressions, then the hashtable contents need
to be recomputed when the Param's value changes.
To avoid unnecessary performance degradation in the case of a Param that
isn't within an aggregate input, add logic to the planner to determine
which Params are within aggregate inputs. This requires a new field in
struct Agg, but fortunately we never write plans to disk, so this isn't
an initdb-forcing change.
Per report from Jeevan Chalke. This has been broken since forever,
so back-patch to all supported branches.
Andrew Gierth, with minor adjustments by me
Report: <CAM2+6=VY8ykfLT5Q8vb9B6EbeBk-NGuLbT6seaQ+Fq4zXvrDcA@mail.gmail.com>
Obvious brain fade in set_rel_consider_parallel(). Noticed it while
adjusting the adjacent RTE_FUNCTION case.
In 9.6, also make the code look more like what I just did in HEAD
by removing the unnecessary function_rte_parallel_ok subroutine
(it does nothing that expression_tree_walker wouldn't do).
We need to scan the whole parse tree for parallel-unsafe functions.
If there are none, we'll later need to determine whether particular
subtrees contain any parallel-restricted functions. The previous coding
retained no knowledge from the first scan, even though this is very
wasteful in the common case where the query contains only parallel-safe
functions. We can bypass all of the later scans by remembering that fact.
This provides a small but measurable speed improvement when the case
applies, and shouldn't cost anything when it doesn't.
Patch by me, reviewed by Robert Haas
Discussion: <3740.1471538387@sss.pgh.pa.us>
We implement a dozen or so parameterless functions that the SQL standard
defines special syntax for. Up to now, that was done by converting them
into more or less ad-hoc constructs such as "'now'::text::date". That's
messy for multiple reasons: it exposes what should be implementation
details to users, and performance is worse than it needs to be in several
cases. To improve matters, invent a new expression node type
SQLValueFunction that can represent any of these parameterless functions.
Bump catversion because this changes stored parsetrees for rules.
Discussion: <30058.1463091294@sss.pgh.pa.us>
ExecEvalCase() tried to save a cycle or two by passing
&econtext->caseValue_isNull as the isNull argument to its sub-evaluation of
the CASE value expression. If that subexpression itself contained a CASE,
then *isNull was an alias for econtext->caseValue_isNull within the
recursive call of ExecEvalCase(), leading to confusion about whether the
inner call's caseValue was null or not. In the worst case this could lead
to a core dump due to dereferencing a null pointer. Fix by not assigning
to the global variable until control comes back from the subexpression.
Also, avoid using the passed-in isNull pointer transiently for evaluation
of WHEN expressions. (Either one of these changes would have been
sufficient to fix the known misbehavior, but it's clear now that each of
these choices was in itself dangerous coding practice and best avoided.
There do not seem to be any similar hazards elsewhere in execQual.c.)
Also, it was possible for inlining of a SQL function that implements the
equality operator used for a CASE comparison to result in one CASE
expression's CaseTestExpr node being inserted inside another CASE
expression. This would certainly result in wrong answers since the
improperly nested CaseTestExpr would be caused to return the inner CASE's
comparison value not the outer's. If the CASE values were of different
data types, a crash might result; moreover such situations could be abused
to allow disclosure of portions of server memory. To fix, teach
inline_function to check for "bare" CaseTestExpr nodes in the arguments of
a function to be inlined, and avoid inlining if there are any.
Heikki Linnakangas, Michael Paquier, Tom Lane
Report: https://github.com/greenplum-db/gpdb/pull/327
Report: <4DDCEEB8.50602@enterprisedb.com>
Security: CVE-2016-5423
Commits 4452000f3 et al established semantics for NullTest.argisrow that
are a bit different from its initial conception: rather than being merely
a cache of whether we've determined the input to have composite type,
the flag now has the further meaning that we should apply field-by-field
testing as per the standard's definition of IS [NOT] NULL. If argisrow
is false and yet the input has composite type, the construct instead has
the semantics of IS [NOT] DISTINCT FROM NULL. Update the comments in
primnodes.h to clarify this, and fix ruleutils.c and deparse.c to print
such cases correctly. In the case of ruleutils.c, this merely results in
cosmetic changes in EXPLAIN output, since the case can't currently arise
in stored rules. However, it represents a live bug for deparse.c, which
would formerly have sent a remote query that had semantics different
from the local behavior. (From the user's standpoint, this means that
testing a remote nested-composite column for null-ness could have had
unexpected recursive behavior much like that fixed in 4452000f3.)
In a related but somewhat independent fix, make plancat.c set argisrow
to false in all NullTest expressions constructed to represent "attnotnull"
constructs. Since attnotnull is actually enforced as a simple null-value
check, this is a more accurate representation of the semantics; we were
previously overpromising what it meant for composite columns, which might
possibly lead to incorrect planner optimizations. (It seems that what the
SQL spec expects a NOT NULL constraint to mean is an IS NOT NULL test, so
arguably we are violating the spec and should fix attnotnull to do the
other thing. If we ever do, this part should get reverted.)
Back-patch, same as the previous commit.
Discussion: <10682.1469566308@sss.pgh.pa.us>
cost_rescan assumed that we don't need to rebuild the hash table when
rescanning a hash join. However, that's currently only true for
single-batch joins; for a multi-batch join we must charge full freight.
This probably has escaped notice because we'd be unlikely to put a hash
join on the inside of a nestloop anyway. Nonetheless, it's wrong.
Fix in HEAD, but don't backpatch for fear of destabilizing plans in
stable releases.
The SQL standard appears to specify that IS [NOT] NULL's tests of field
nullness are non-recursive, ie, we shouldn't consider that a composite
field with value ROW(NULL,NULL) is null for this purpose.
ExecEvalNullTest got this right, but eval_const_expressions did not,
leading to weird inconsistencies depending on whether the expression
was such that the planner could apply constant folding.
Also, adjust the docs to mention that IS [NOT] DISTINCT FROM NULL can be
used as a substitute test if a simple null check is wanted for a rowtype
argument. That motivated reordering things so that IS [NOT] DISTINCT FROM
is described before IS [NOT] NULL. In HEAD, I went a bit further and added
a table showing all the comparison-related predicates.
Per bug #14235. Back-patch to all supported branches, since it's certainly
undesirable that constant-folding should change the semantics.
Report and patch by Andrew Gierth; assorted wordsmithing and revised
regression test cases by me.
Report: <20160708024746.1410.57282@wrigleys.postgresql.org>
The aggfilter expression should be removed from the parent (combining)
Aggref, since it's not supposed to apply the filter, and indeed cannot
because any Vars used in the filter would not be available after the
lower-level aggregation step. Per report from Jeff Janes.
(This has been broken since the introduction of partial aggregation,
I think. The error became obvious after commit 59a3795c2, when setrefs.c
began processing the parent Aggref's fields normally and thus would detect
such Vars. The special-case coding previously used in setrefs.c skipped
over the parent's aggfilter field without processing it. That was broken
in its own way because no other setrefs.c processing got applied either;
though since the executor would not execute the filter expression, only
initialize it, that oversight might not have had any visible symptoms at
present.)
Report: <CAMkU=1xfuPf2edAe4ZGXTmJpU7jxuKukKyvNtEXwu35B7dvejg@mail.gmail.com>
We must not push down a foreign join when the foreign tables involved
should be accessed under different user mappings. Previously we tried
to enforce that rule literally during planning, but that meant that the
resulting plans were dependent on the current contents of the
pg_user_mapping catalog, and we had to blow away all cached plans
containing any remote join when anything at all changed in pg_user_mapping.
This could have been improved somewhat, but the fact that a syscache inval
callback has very limited info about what changed made it hard to do better
within that design. Instead, let's change the planner to not consider user
mappings per se, but to allow a foreign join if both RTEs have the same
checkAsUser value. If they do, then they necessarily will use the same
user mapping at runtime, and we don't need to know specifically which one
that is. Post-plan-time changes in pg_user_mapping no longer require any
plan invalidation.
This rule does give up some optimization ability, to wit where two foreign
table references come from views with different owners or one's from a view
and one's directly in the query, but nonetheless the same user mapping
would have applied. We'll sacrifice the first case, but to not regress
more than we have to in the second case, allow a foreign join involving
both zero and nonzero checkAsUser values if the nonzero one is the same as
the prevailing effective userID. In that case, mark the plan as only
runnable by that userID.
The plancache code already had a notion of plans being userID-specific,
in order to support RLS. It was a little confused though, in particular
lacking clarity of thought as to whether it was the rewritten query or just
the finished plan that's dependent on the userID. Rearrange that code so
that it's clearer what depends on which, and so that the same logic applies
to both RLS-injected role dependency and foreign-join-injected role
dependency.
Note that this patch doesn't remove the other issue mentioned in the
original complaint, which is that while we'll reliably stop using a foreign
join if it's disallowed in a new context, we might fail to start using a
foreign join if it's now allowed, but we previously created a generic
cached plan that didn't use one. It was agreed that the chance of winning
that way was not high enough to justify the much larger number of plan
invalidations that would have to occur if we tried to cause it to happen.
In passing, clean up randomly-varying spelling of EXPLAIN commands in
postgres_fdw.sql, and fix a COSTS ON example that had been allowed to
leak into the committed tests.
This reverts most of commits fbe5a3fb7 and 5d4171d1c, which were the
previous attempt at ensuring we wouldn't push down foreign joins that
span permissions contexts.
Etsuro Fujita and Tom Lane
Discussion: <d49c1e5b-f059-20f4-c132-e9752ee0113e@lab.ntt.co.jp>
Test the external-sort code path in CLUSTER for two different scenarios:
multiple-pass external sorting, and the best case for replacement
selection, where only one run is produced, so that no merge is required.
This test would have caught the bug fixed in commit 1b0fc8507, at
least when run with valgrind enabled.
In passing, add a short-circuit test in plan_cluster_use_sort() to make
dead certain that it selects sorting when enable_indexscan is off. As
things stand, that would happen anyway, but it seems like good future
proofing for this test.
Peter Geoghegan
Discussion: <CAM3SWZSgxehDkDMq1FdiW2A0Dxc79wH0hz1x-TnGy=1BXEL+nw@mail.gmail.com>
There isn't really any reason not to; the original comments here were
partly confused about subplans versus subquery-in-FROM, and partly
dependent on restrictions that no longer apply now that subqueries return
Paths not Plans. Depending on what's inside the subquery, it might fail
to produce any parallel_safe Paths, but that's fine.
Tom Lane and Robert Haas
The previous coding assumed that the value derived by
set_rel_consider_parallel() for an appendrel parent would be accurate for
all the appendrel's children; but this is not so, for example because one
child might scan a temp table. Instead, apply set_rel_consider_parallel()
to each child rel as well as the parent, and then take the AND of the
results as controlling parallel safety for the appendrel as a whole.
(We might someday be able to deal more intelligently than this with cases
in which some of the childrels are parallel-safe and others not, but that's
for later.)
Robert Haas and Tom Lane
This was the intention all along, but an extraneous "return;" in
set_rel_consider_parallel() caused sampled rels to never be marked
consider_parallel.
Since we don't have any partial tablesample path/plan type yet, there's
no possibility of parallelizing the sample scan itself; but this fix
allows such a scan to appear below a parallel join, for example.
We were just leaving the cost fields zeroes, which produces obviously bogus
output with force_parallel_mode = on. With force_parallel_mode = regress,
the zeroes are hidden, but I wonder if they wouldn't still confuse add-on
code such as auto_explain.
I'd been wondering why I was sometimes seeing fractional rowcount
estimates in parallel-query situations, and this seems to be the
reason. (You won't see the fractional parts in EXPLAIN, because it
prints rowcounts with %.0f, but they are apparent in the debugger.)
A fractional rowcount is not any saner for a partial path than any
other kind of path, and it's equally likely to break cost estimation
for higher paths, so apply clamp_row_est() like we do in other places.
In commit 915b703e1 I gave get_agg_clause_costs() the responsibility of
marking Aggref nodes with the appropriate aggtranstype. I failed to notice
that where it was being called from, it might see only a subset of the
Aggref nodes that were in the original targetlist. Specifically, if there
are duplicate aggregate calls in the tlist, either make_sort_input_target
or make_window_input_target might put just a single instance into the
grouping_target, and then only that instance would get marked. Fix by
moving the call back into grouping_planner(), before we start building
assorted PathTargets from the query tlist. Per report from Stefan Huehner.
Report: <20160702131056.GD3165@huehner.biz>
In commit 68fa28f77 I tried to teach SS_finalize_plan() to cope with
initPlans attached anywhere in the plan tree, by dint of moving its
handling of those into the recursion in finalize_plan(). It turns out that
that doesn't really work: if a lower-level plan node emits an initPlan
output parameter in its targetlist, it's legitimate for upper levels to
reference those Params --- and at the point where this code runs, those
references look just like the Param itself, so finalize_plan() quite
properly rejects them as being in the wrong place. We could lobotomize
the checks enough to allow that, probably, but then it's not clear that
we'd have any meaningful check for misplaced Params at all. What seems
better, at least in the near term, is to tweak standard_planner() a bit
so that initPlans are never placed anywhere but the topmost plan node
for a query level, restoring the behavior that occurred pre-9.6. Possibly
we can do better if this code is ever merged into setrefs.c: then it would
be possible to check a Param's placement only when we'd failed to replace
it with a Var referencing a child plan node's targetlist.
BTW, I'm now suspicious that finalize_plan is doing the wrong thing by
returning the node's allParam rather than extParam to be incorporated
in the parent node's set of used parameters. However, it makes no
difference given that initPlans only appear at top level, so I'll leave
that alone for now.
Another thing that emerged from this is that standard_planner() needs
to check for initPlans before deciding that it's safe to stick a Gather
node on top in force_parallel_mode mode. We previously guarded against
that by deciding the plan wasn't wholePlanParallelSafe if any subplans
had been found, but after commit 5ce5e4a12 it's necessary to have this
substitute test, because path parallel_safe markings don't account for
initPlans. (Normally, we'd have decided the paths weren't safe anyway
due to appearances of SubPlan nodes, Params, or CTE scans somewhere in
the tree --- but it's possible for those all to be optimized away while
initPlans still remain.)
Per fuzz testing by Andreas Seltenreich.
Report: <874m89rw7x.fsf@credativ.de>
In the previous design, the GetForeignUpperPaths FDW callback hook was
called before we got around to labeling upper relations with the proper
consider_parallel flag; this meant that any upper paths created by an FDW
would be marked not-parallel-safe. While that's probably just as well
right now, we aren't going to want it to be true forever. Hence, abandon
the idea that FDWs should be allowed to inject upper paths before the core
code has gotten around to creating the relevant upper relation. (Well,
actually they still can, but it's on their own heads how well it works.)
Instead, adopt the same API already designed for create_upper_paths_hook:
we call GetForeignUpperPaths after each upperrel has been created and
populated with the paths the core planner knows how to make.
Commit 3fc6e2d7f5 introduced new "upper"
RelOptInfo structures but didn't set consider_parallel for them
correctly, a point I completely missed when reviewing it. Later,
commit e06a38965b made the situation
worse by doing it incorrectly for the grouping relation. Try to
straighten all of that out. Along the way, get rid of the annoying
wholePlanParallelSafe flag, which was only necessarily because of
the fact that upper planning stages didn't use paths at the time
that code was written.
The most important immediate impact of these changes is that
force_parallel_mode will provide useful test coverage in quite a few
more scenarios than it did previously, but it's also necessary
preparation for fixing some problems related to subqueries.
Patch by me, reviewed by Tom Lane.
VS2013 apparently has a problem with taking the address of a formal
parameter in some cases. We do that elsewhere without trouble, but
in this case the address is being passed to a subroutine that will
probably get inlined, so maybe the combination of those things is
what tickles the bug. Anyway, introducing an extra copy of the
parameter value is enough to work around it. Per trouble report
from Umair Shahid.
Report: <CAM184AcjqKYZSdQqBHDrnENXHhW=mXbUC46QYPJ=nAh0gUHCGA@mail.gmail.com>
Since get_relation_foreign_keys doesn't try to determine whether RTEs
are actually part of the query semantics, it might make FK info records
linking to RTEs that won't have a RelOptInfo at all. Cope with that.
Per bug #14219 from Andrew Gierth.
Report: <20160629183338.1397.43514@wrigleys.postgresql.org>
If we need to use a gating Result node for pseudoconstant quals,
create_scan_plan() intentionally suppresses use_physical_tlist's checks
on whether there are matches for sortgroupref labels, on the grounds that
we don't need matches because we can label the Result's projection output
properly. However, it then called apply_pathtarget_labeling_to_tlist
anyway. This oversight was harmless when written, but in commit aeb9ae645
I made that function throw an error if there was no match. Thus, the
combination of a table scan, pseudoconstant quals, and a non-simple-Var
sortgroupref column threw the dreaded "ORDER/GROUP BY expression not found
in targetlist" error. To fix, just skip applying the labeling in this
case. Per report from Rushabh Lathia.
Report: <CAGPqQf2iLB8t6t-XrL-zR233DFTXxEsfVZ4WSqaYfLupEoDxXA@mail.gmail.com>
It's rather silly to make a separate pass over the tlist + HAVING qual,
and a separate set of visits to the syscache, when get_agg_clause_costs
already has all the required information in hand. This nets out as less
code as well as fewer cycles.
The original coding had three separate booleans representing partial
aggregation behavior, which was confusing, unreadable, and error-prone,
not least because the booleans weren't always listed in the same order.
It was also inadequate for the allegedly-desirable future extension to
support intermediate partial aggregation, because we'd need separate
markers for serialization and deserialization in such a case.
Merge these bools into an enum "AggSplit" to provide symbolic names for
the supported operating modes (and document what those are). By assigning
the values of the enum constants carefully, we can treat AggSplit values
as options bitmasks so that tests of what to do aren't noticeably more
expensive than before.
While at it, get rid of Aggref.aggoutputtype. That's not needed since
commit 59a3795c2 got rid of setrefs.c's special-purpose Aggref comparison
code, and it likewise seemed more confusing than helpful.
Assorted comment cleanup as well (there's still more that I want to do
in that line).
catversion bump for change in Aggref node contents. Should be the last
one for partial-aggregation changes.
Discussion: <29309.1466699160@sss.pgh.pa.us>
Commit e06a38965's original coding for constructing the execution-time
expression tree for a combining aggregate was rather messy, involving
duplicating quite a lot of code in setrefs.c so that it could inject
a nonstandard matching rule for Aggrefs. Get rid of that in favor of
explicitly constructing a combining Aggref with a partial Aggref as input,
then allowing setref's normal matching logic to match the partial Aggref
to the output of the lower plan node and hence replace it with a Var.
In passing, rename and redocument make_partialgroup_input_target to have
some connection to what it actually does.
The original specification for this called for the deserialization function
to have signature "deserialize(serialtype) returns transtype", which is a
security violation if transtype is INTERNAL (which it always would be in
practice) and serialtype is not (which ditto). The patch blithely overrode
the opr_sanity check for that, which was sloppy-enough work in itself,
but the indisputable reason this cannot be allowed to stand is that CREATE
FUNCTION will reject such a signature and thus it'd be impossible for
extensions to create parallelizable aggregates.
The minimum fix to make the signature type-safe is to add a second, dummy
argument of type INTERNAL. But to lock it down a bit more and make misuse
of INTERNAL-accepting functions less likely, let's get rid of the ability
to specify a "serialtype" for an aggregate and just say that the only
useful serialtype is BYTEA --- which, in practice, is the only interesting
value anyway, due to the usefulness of the send/recv infrastructure for
this purpose. That means we only have to allow "serialize(internal)
returns bytea" and "deserialize(bytea, internal) returns internal" as
the signatures for these support functions.
In passing fix bogus signature of int4_avg_combine, which I found thanks
to adding an opr_sanity check on combinefunc signatures.
catversion bump due to removing pg_aggregate.aggserialtype and adjusting
signatures of assorted built-in functions.
David Rowley and Tom Lane
Discussion: <27247.1466185504@sss.pgh.pa.us>
The original upper-planner-pathification design (commit 3fc6e2d7f5)
assumed that we could always determine during Path formation whether or not
we would need a Result plan node to perform projection of a targetlist.
That turns out not to work very well, though, because createplan.c still
has some responsibilities for choosing the specific target list associated
with sorting/grouping nodes (in particular it might choose to add resjunk
columns for sorting). We might not ever refactor that --- doing so would
push more work into Path formation, which isn't attractive --- and we
certainly won't do so for 9.6. So, while create_projection_path and
apply_projection_to_path can tell for sure what will happen if the subpath
is projection-capable, they can't tell for sure when it isn't. This is at
least a latent bug in apply_projection_to_path, which might think it can
apply a target to a non-projecting node when the node will end up computing
something different.
Also, I'd tied the creation of a ProjectionPath node to whether or not a
Result is needed, but it turns out that we sometimes need a ProjectionPath
node anyway to avoid modifying a possibly-shared subpath node. Callers had
to use create_projection_path for such cases, and we added code to them
that knew about the potential omission of a Result node and attempted to
adjust the cost estimates for that. That was uncertainly correct and
definitely ugly/unmaintainable.
To fix, have create_projection_path explicitly check whether a Result
is needed and adjust its cost estimate accordingly, though it creates
a ProjectionPath in either case. apply_projection_to_path is now mostly
just an optimized version that can avoid creating an extra Path node when
the input is known to not be shared with any other live path. (There
is one case that create_projection_path doesn't handle, which is pushing
parallel-safe expressions below a Gather node. We could make it do that
by duplicating the GatherPath, but there seems no need as yet.)
create_projection_plan still has to recheck the tlist-match condition,
which means that if the matching situation does get changed by createplan.c
then we'll have made a slightly incorrect cost estimate. But there seems
no help for that in the near term, and I doubt it occurs often enough,
let alone would change planning decisions often enough, to be worth
stressing about.
I added a "dummypp" field to ProjectionPath to track whether
create_projection_path thinks a Result is needed. This is not really
necessary as-committed because create_projection_plan doesn't look at the
flag; but it seems like a good idea to remember what we thought when
forming the cost estimate, if only for debugging purposes.
In passing, get rid of the target_parallel parameter added to
apply_projection_to_path by commit 54f5c5150. I don't think that's a good
idea because it involves callers in what should be an internal decision,
and opens us up to missing optimization opportunities if callers think they
don't need to provide a valid flag, as most don't. For the moment, this
just costs us an extra has_parallel_hazard call when planning a Gather.
If that starts to look expensive, I think a better solution would be to
teach PathTarget to carry/cache knowledge of parallel-safety of its
contents.
This patch provides a new implementation of the logic added by commit
137805f89 and later removed by 77ba61080. It differs from the original
primarily in expending much less effort per joinrel in large queries,
which it accomplishes by doing most of the matching work once per query not
once per joinrel. Hopefully, it's also less buggy and better commented.
The never-documented enable_fkey_estimates GUC remains gone.
There remains work to be done to make the selectivity estimates account
for nulls in FK referencing columns; but that was true of the original
patch as well. We may be able to address this point later in beta.
In the meantime, any error should be in the direction of overestimating
rather than underestimating joinrel sizes, which seems like the direction
we want to err in.
Tomas Vondra and Tom Lane
Discussion: <31041.1465069446@sss.pgh.pa.us>
The previous code neglected the fact that the scanjoin_target might
carry sortgroupref labelings that we need to absorb. Instead, do
create_projection_path() unconditionally, and tweak the path's cost
estimate after the fact. (I'm now convinced that we ought to refactor
the way we account for sometimes not needing a separate projection step,
but right now is not the time for that sort of cleanup.)
Problem identified by Amit Kapila, patch by me.
When doing partial aggregation, the args list of the upper (combining)
Aggref node is replaced by a Var representing the output of the partial
aggregation steps, which has either the aggregate's transition data type
or a serialized representation of that. However, nodeAgg.c blindly
continued to use the args list as an indication of the user-level argument
types. This broke resolution of polymorphic transition datatypes at
executor startup (though it accidentally failed to fail for the ANYARRAY
case, which is likely the only one anyone had tested). Moreover, the
constructed FuncExpr passed to the finalfunc contained completely wrong
information, which would have led to bogus answers or crashes for any case
where the finalfunc examined that information (which is only likely to be
with polymorphic aggregates using a non-polymorphic transition type).
As an independent bug, apply_partialaggref_adjustment neglected to resolve
a polymorphic transition datatype before assigning it as the output type
of the lower-level Aggref node. This again accidentally failed to fail
for ANYARRAY but would be unlikely to work in other cases.
To fix the first problem, record the user-level argument types in a
separate OID-list field of Aggref, and look to that rather than the args
list when asking what the argument types were. (It turns out to be
convenient to include any "direct" arguments in this list too, although
those are not currently subject to being overwritten.)
Rather than adding yet another resolve_aggregate_transtype() call to fix
the second problem, add an aggtranstype field to Aggref, and store the
resolved transition type OID there when the planner first computes it.
(By doing this in the planner and not the parser, we can allow the
aggregate's transition type to change from time to time, although no DDL
support yet exists for that.) This saves nothing of consequence for
simple non-polymorphic aggregates, but for polymorphic transition types
we save a catalog lookup during executor startup as well as several
planner lookups that are new in 9.6 due to parallel query planning.
In passing, fix an error that was introduced into count_agg_clauses_walker
some time ago: it was applying exprTypmod() to something that wasn't an
expression node at all, but a TargetEntry. exprTypmod silently returned
-1 so that there was not an obvious failure, but this broke the intended
sensitivity of aggregate space consumption estimates to the typmod of
varchar and similar data types. This part needs to be back-patched.
Catversion bump due to change of stored Aggref nodes.
Discussion: <8229.1466109074@sss.pgh.pa.us>
Commit 04ae11f62e removed some broken
code to apply the scan/join target to partial paths, but its theory
that this processing step is totally unnecessary turns out to be wrong.
Put similar code back again, but this time, check for parallel-safety
and avoid in-place modifications to paths that may already have been
used as part of some other path.
(This is not an entirely elegant solution to this problem; it might
be better, for example, to postpone generate_gather_paths for the
topmost scan/join rel until after the scan/join target has been
applied. But this is not the time for such redesign work.)
Amit Kapila and Robert Haas
The main point of doing this is to allow the cutoff to be set very small,
even zero, to allow parallel-query behavior to be tested on relatively
small tables such as we typically use in the regression tests. But it
might be of use to users too. The number-of-workers scaling behavior in
create_plain_partial_paths() is pretty ad-hoc and subject to change, so
we won't expose anything about that, but the notion of not considering
parallel query at all for tables below size X seems reasonably stable.
Amit Kapila, per a suggestion from me
Discussion: <17170.1465830165@sss.pgh.pa.us>
The struct definition for PathTarget specifies that a NULL sortgrouprefs
pointer means no sortgroupref labels. While it's likely that there
should always be at least one labeled column in the places that were
unconditionally fetching through the pointer, it seems wiser to adhere to
the data structure specification and test first. Add a macro to make this
convenient. Per experimentation with running the regression tests with a
very small parallelization threshold --- the crash I observed may well
represent a bug elsewhere, but still this coding was not very robust.
Report: <20756.1465834072@sss.pgh.pa.us>
Commit b12fd41c6 added a "reltarget_has_non_vars" field to RelOptInfo,
but failed to maintain it accurately. Since its only purpose was to skip
calls to has_parallel_hazard() in the simple case where a rel's targetlist
is all Vars, and that call is really pretty cheap in that case anyway, it
seems like this is just a case of premature optimization. Let's drop the
flag and do the calls unconditionally until it's proven that we need more
smarts here.
As noted by Andres Freund, we'd accumulated quite a few similar functions
in clauses.c that examine all functions in an expression tree to see if
they satisfy some boolean test. Reduce the duplication by inventing a
function check_functions_in_node() that applies a simple callback function
to each SQL function OID appearing in a given expression node. This also
fixes some arguable oversights; for example, contain_mutable_functions()
did not check aggregate or window functions for mutability. I doubt that
that represents a live bug at the moment, because we don't really consider
mutability for aggregates; but it might someday be one.
I chose to put check_functions_in_node() in nodeFuncs.c because it seemed
like other modules might wish to use it in future. That in turn forced
moving set_opfuncid() et al into nodeFuncs.c, as the alternative was for
nodeFuncs.c to depend on optimizer/setrefs.c which didn't seem very clean.
In passing, teach contain_leaked_vars_walker() about a few more expression
node types it can safely look through, and improve the rather messy and
undercommented code in has_parallel_hazard_walker().
Discussion: <20160527185853.ziol2os2zskahl7v@alap3.anarazel.de>
Transmit the leader's temp-namespace state to workers. This is important
because without it, the workers do not really have the same search path
as the leader. For example, there is no good reason (and no extant code
either) to prevent a worker from executing a temp function that the
leader created previously; but as things stood it would fail to find the
temp function, and then either fail or execute the wrong function entirely.
We still prohibit a worker from creating a temp namespace on its own.
In effect, a worker can only see the session's temp namespace if the leader
had created it before starting the worker, which seems like the right
semantics.
Also, transmit the leader's BackendId to workers, and arrange for workers
to use that when determining the physical file path of a temp relation
belonging to their session. While the original intent was to prevent such
accesses entirely, there were a number of holes in that, notably in places
like dbsize.c which assume they can safely access temp rels of other
sessions anyway. We might as well get this right, as a small down payment
on someday allowing workers to access the leader's temp tables. (With
this change, directly using "MyBackendId" as a relation or buffer backend
ID is deprecated; you should use BackendIdForTempRelations() instead.
I left a couple of such uses alone though, as they're not going to be
reachable in parallel workers until we do something about localbuf.c.)
Move the thou-shalt-not-access-thy-leader's-temp-tables prohibition down
into localbuf.c, which is where it actually matters, instead of having it
in relation_open(). This amounts to recognizing that access to temp
tables' catalog entries is perfectly safe in a worker, it's only the data
in local buffers that is problematic.
Having done all that, we can get rid of the test in has_parallel_hazard()
that says that use of a temp table's rowtype is unsafe in parallel workers.
That test was unduly expensive, and if we really did need such a
prohibition, that was not even close to being a bulletproof guard for it.
(For example, any user-defined function executed in a parallel worker
might have attempted such access.)
Such paths are unsafe. To make it cheaper to detect when this case
applies, track whether a relation's default PathTarget contains any
non-Vars. In most cases, the answer will be no, which enables us to
determine cheaply that the target list for a proposed path is
parallel-safe. However, subquery pull-up can create cases that
require us to inspect the target list more carefully.
Amit Kapila, reviewed by me.
Fix a couple of overlooked uses of "degree" terminology. Make the parallel
worker count selection logic in create_plain_partial_paths more robust (in
particular, it failed with max_parallel_workers_per_gather set to zero).
This terminology provoked widespread complaints. So, instead, rename
the GUC max_parallel_degree to max_parallel_workers_per_gather
(leaving room for a possible future GUC max_parallel_workers that acts
as a system-wide limit), and rename the parallel_degree reloption to
parallel_workers. Rename structure members to match.
These changes create a dump/restore hazard for users of PostgreSQL
9.6beta1 who have set the reloption (or applied the GUC using ALTER
USER or ALTER DATABASE).
This commit reverts 137805f89 as well as the associated commits 015e88942,
5306df283, and 68d704edb. We found multiple bugs in this feature, and
there was concern about possible planner slowdown (though to be fair,
exhibiting a very large slowdown proved difficult). The way forward
requires a considerable rewrite, which may or may not be possible to
accomplish in time for beta2. In my judgment reviewing the rewrite will
be easier to accomplish starting from a clean slate, so let's temporarily
revert what's there now. This also leaves us in a safe state if it turns
out to be necessary to postpone the rewrite to the next development cycle.
Discussion: <20160429102531.GA13701@huehner.biz>
The partial paths that get modified may already have been used as
part of a GatherPath which appears in the path list, so modifying
them is not a good idea at this stage - especially because this
code has no check that the PathTarget is in fact parallel-safe.
When partial aggregation is being performed, this is actually
harmless because we'll end up replacing the pathtargets here with
the correct ones within create_grouping_paths(). But if we've got
a query tree containing only scan/join operations then this can
result in incorrectly pushing down parallel-restricted target
list entries. If those are, for example, references to subqueries,
that can crash the server; but it's wrong in any event.
Amit Kapila
Mostly these are just comments but there are a few in documentation
and a handful in code and tests. Hopefully this doesn't cause too much
unnecessary pain for backpatching. I relented from some of the most
common like "thru" for that reason. The rest don't seem numerous
enough to cause problems.
Thanks to Kevin Lyda's tool https://pypi.python.org/pypi/misspellings
As part of upper planner pathification (commit 3fc6e2d7f5) I redid
createplan.c's approach to the physical-tlist optimization, in which scan
nodes are allowed to return exactly the underlying table's columns so as
to save doing a projection step at runtime. The logic was intentionally
more aggressive than before about applying the optimization, which is
generally a good thing, but Andres Freund found a case in which it got
too aggressive. Namely, if any column is referenced more than once in
the parent plan node's sorting or grouping column list, we can't optimize
because then that column would need to have more than one ressortgroupref
label, and we only have space for one.
Add logic to detect this situation in use_physical_tlist(), and also add
some error checking in apply_pathtarget_labeling_to_tlist(), which this
example proves was being overly cavalier about whether what it was doing
made any sense.
The added test case exposes the problem only because we do not eliminate
duplicate grouping keys. That might be something to fix someday, but it
doesn't seem like appropriate post-beta work.
Report: <20160526021235.w4nq7k3gnheg7vit@alap3.anarazel.de>
While it could be argued that rejecting system column mentions in the
ON CONFLICT list is an unsupported feature, falling over altogether
just because the table has a unique index on OID is indubitably a bug.
As far as I can tell, fixing infer_arbiter_indexes() is sufficient to
make ON CONFLICT (oid) actually work, though making a regression test
for that case is problematic because of the impossibility of setting
the OID counter to a known value.
Minor cosmetic cleanups along with the bug fix.
subquery_planner() failed to apply expression preprocessing to the
arbiterElems and arbiterWhere fields of an OnConflictExpr. No doubt the
theory was that this wasn't necessary because we don't actually try to
execute those expressions; but that's wrong, because it results in failure
to match to index expressions or index predicates that are changed at all
by preprocessing. Per bug #14132 from Reynold Smith.
Also add pullup_replace_vars processing for onConflictWhere. Perhaps
it's impossible to have a subquery reference there, but I'm not exactly
convinced; and even if true today it's a failure waiting to happen.
Also add some comments to other places where one or another field of
OnConflictExpr is intentionally ignored, with explanation as to why it's
okay to do so.
Also, catalog/dependency.c failed to record any dependency on the named
constraint in ON CONFLICT ON CONSTRAINT, allowing such a constraint to
be dropped while rules exist that depend on it, and allowing pg_dump to
dump such a rule before the constraint it refers to. The normal execution
path managed to error out reasonably for a dangling constraint reference,
but ruleutils.c dumped core; so in addition to fixing the omission, add
a protective check in ruleutils.c, since we can't retroactively add a
dependency in existing databases.
Back-patch to 9.5 where this code was introduced.
Report: <20160510190350.2608.48667@wrigleys.postgresql.org>
Discussion is still underway as to whether to revert the entire patch
that added this function, but that discussion may not conclude before
beta1. So, in the meantime, let's do at least this much.
David Rowley
Now that Paths have their own rows field, print that rather than
the parent relation's rowcount.
Show the relid sets associated with Paths using table names rather
than numbers; since this code is able to print simple Var references
using table names, it seems a bit silly that print_relids can't.
Print the cheapest_parameterized_paths list for a RelOptInfo, and
include information about a parameterized path's required_outer rels.
Noted while trying to use this feature to debug Alexander Kirkouski's
recent bug report.
We mustn't run generate_gather_paths() during add_paths_to_joinrel(),
because that function can be invoked multiple times for the same target
joinrel. Not only is it wasteful to build GatherPaths repeatedly, but
a later add_partial_path() could delete the partial path that a previously
created GatherPath depends on. Instead establish the convention that we
do generate_gather_paths() for a rel only just before set_cheapest().
The code was accidentally not broken for baserels, because as of today there
never is more than one partial path for a baserel. But that assumption
obviously has a pretty short half-life, so move the generate_gather_paths()
calls for those cases as well.
Also add some generic comments explaining how and why this all works.
Per fuzz testing by Andreas Seltenreich.
Report: <871t5pgwdt.fsf@credativ.de>
Given a three-or-more-way equivalence class, such as X.Y = Y.Y = Z.Z,
it was possible for the planner to omit one of the quals needed to
enforce that all members of the equivalence class are actually equal.
This only happened in the case of a parameterized join node for two
of the relations, that is a plan tree like
Nested Loop
-> Scan X
-> Nested Loop
-> Scan Y
-> Scan Z
Filter: Z.Z = X.X
The eclass machinery normally expects to apply X.X = Y.Y when those
two relations are joined, but in this shape of plan tree they aren't
joined until the top node --- and, if the lower nested loop is marked
as parameterized by X, the top node will assume that the relevant eclass
condition(s) got pushed down into the lower node. On the other hand,
the scan of Z assumes that it's only responsible for constraining Z.Z
to match any one of the other eclass members. So one or another of
the required quals sometimes fell between the cracks, depending on
whether consideration of the eclass in get_joinrel_parampathinfo()
for the lower nested loop chanced to generate X.X = Y.Y or X.X = Z.Z
as the appropriate constraint there. If it generated the latter,
it'd erroneously suppose that the Z scan would take care of matters.
To fix, force X.X = Y.Y to be generated and applied at that join node
when this case occurs.
This is *extremely* hard to hit in practice, because various planner
behaviors conspire to mask the problem; starting with the fact that the
planner doesn't really like to generate a parameterized plan of the
above shape. (It might have been impossible to hit it before we
tweaked things to allow this plan shape for star-schema cases.) Many
thanks to Alexander Kirkouski for submitting a reproducible test case.
The bug can be demonstrated in all branches back to 9.2 where parameterized
paths were introduced, so back-patch that far.
The way that PartialAggregate and FinalizeAggregate plan nodes were
displaying output columns before was bogus. Now, FinalizeAggregate
produces the same outputs as an Aggregate would have produced, while
PartialAggregate produces each of those outputs prefixed by the word
PARTIAL.
Discussion: 12585.1460737650@sss.pgh.pa.us
Patch by me, reviewed by David Rowley.
Change max_parallel_degree default from 0 to 2. It is possible that
this is not a good idea, or that we should go with 1 worker rather
than 2, but we won't find out without trying it. Along the way,
reword the documentation for max_parallel_degree a little bit to
hopefully make it more clear.
Discussion: 20160420174631.3qjjhpwsvvx5bau5@alap3.anarazel.de
Given a left join containing a full join in its righthand side, with
the left join's joinclause referencing only one side of the full join
(in a non-strict fashion, so that the full join doesn't get simplified),
the planner could fail with "failed to build any N-way joins" or related
errors. This happened because the full join was seen as overlapping the
left join's RHS, and then recent changes within join_is_legal() caused
that function to conclude that the full join couldn't validly be formed.
Rather than try to rejigger join_is_legal() yet more to allow this,
I think it's better to fix initsplan.c so that the required join order
is explicit in the SpecialJoinInfo data structure. The previous coding
there essentially ignored full joins, relying on the fact that we don't
flatten them in the joinlist data structure to preserve their ordering.
That's sufficient to prevent a wrong plan from being formed, but as this
example shows, it's not sufficient to ensure that the right plan will
be formed. We need to work a bit harder to ensure that the right plan
looks sane according to the SpecialJoinInfos.
Per bug #14105 from Vojtech Rylko. This was apparently induced by
commit 8703059c6 (though now that I've seen it, I wonder whether there
are related cases that could have failed before that); so back-patch
to all active branches. Unfortunately, that patch also went into 9.0,
so this bug is a regression that won't be fixed in that branch.
The original patch kind of ignored the fact that we were doing something
different from a costing point of view, but nobody noticed. This patch
fixes that oversight.
David Rowley
Per discussion, this gives potential users of the hook more flexibility,
because they can build custom Paths that implement only one stage of
upper processing atop core-provided Paths for earlier stages.
Coverity complained that the code added by 015e88942a lacked an
error check for SearchSysCache1 failures, which it should have. But
the code was pretty duff in other ways too, including failure to think
about whether it could really cope with arrays of different lengths.
Now indexes (but only B-tree for now) can contain "extra" column(s) which
doesn't participate in index structure, they are just stored in leaf
tuples. It allows to use index only scan by using single index instead
of two or more indexes.
Author: Anastasia Lubennikova with minor editorializing by me
Reviewers: David Rowley, Peter Geoghegan, Jeff Janes
The code that estimates what parallel degree should be uesd for the
scan of a relation is currently rather stupid, so add a parallel_degree
reloption that can be used to override the planner's rather limited
judgement.
Julien Rouhaud, reviewed by David Rowley, James Sewell, Amit Kapila,
and me. Some further hacking by me.
We still use replacement selection for the first run of the sort only
and only when the number of tuples is relatively small. Otherwise,
the first run, and subsequent runs in all cases, are produced using
quicksort. This tends to be faster except perhaps for very small
amounts of working memory.
Peter Geoghegan, reviewed by Tomas Vondra, Jeff Janes, Mithun Cy,
Greg Stark, and me.
In cases where joins use multiple columns we currently assess each join
separately causing gross mis-estimates for join cardinality.
This patch adds use of FK information for the first time into the
planner. When FKs are present and we have multi-column join information,
plan estimates will be drastically improved. Cases with multiple FKs
are handled, though partial matches are ignored currently.
Net effect is substantial performance improvements for joins in many
common cases. Additional planning time is isolated to cases that are
currently performing poorly, measured at 0.08 - 0.15 ms.
Please watch for planner performance regressions; circumstances seem
unlikely but the law of unintended consequences may apply somewhen.
Additional complex tests welcome to prove this before release.
Tests can be performed using SET enable_fkey_estimates = on | off
using scripts provided during Hackers discussions, message id:
552335D9.3090707@2ndquadrant.com
Authors: Tomas Vondra and David Rowley
Reviewed and tested by Simon Riggs, adding comments only
Extracted from pending unique-join patch, since this is a rather large
delta but it's simply moving code out into separately-accessible
subroutines.
I (tgl) did choose to add a bit more logic to rel_supports_distinctness,
so that it verifies that there's at least one potentially usable unique
index rather than just checking indexlist != NIL. Otherwise there's
no functional change here.
David Rowley
has_parallel_hazard() was ignoring the proparallel markings for
aggregates, which is no good. Fix that. There was no way to mark
an aggregate as actually being parallel-safe, either, so add a
PARALLEL option to CREATE AGGREGATE.
Patch by me, reviewed by David Rowley.
Previously, the planner would reject an index-only scan if any restriction
clause for its table used a column not available from the index, even
if that restriction clause would later be dropped from the plan entirely
because it's implied by the index's predicate. This is a fairly common
situation for partial indexes because predicates using columns not included
in the index are often the most useful kind of predicate, and we have to
duplicate (or at least imply) the predicate in the WHERE clause in order
to get the index to be considered at all. So index-only scans were
essentially unavailable with such partial indexes.
To fix, we have to do detection of implied-by-predicate clauses much
earlier in the planner. This patch puts it in check_index_predicates
(nee check_partial_indexes), meaning it gets done for every partial index,
whereas we previously only considered this issue at createplan time,
so that the work was only done for an index actually selected for use.
That could result in a noticeable planning slowdown for queries against
tables with many partial indexes. However, testing suggested that there
isn't really a significant cost, especially not with reasonable numbers
of partial indexes. We do get a small additional benefit, which is that
cost_index is more accurate since it correctly discounts the evaluation
cost of clauses that will be removed. We can also avoid considering such
clauses as potential indexquals, which saves useless matching cycles in
the case where the predicate columns aren't in the index, and prevents
generating bogus plans that double-count the clause's selectivity when
the columns are in the index.
Tomas Vondra and Kyotaro Horiguchi, reviewed by Kevin Grittner and
Konstantin Knizhnik, and whacked around a little by me
This is necessary infrastructure for supporting parallel aggregation
for aggregates whose transition type is "internal". Such values
can't be passed between cooperating processes, because they are
just pointers.
David Rowley, reviewed by Tomas Vondra and by me.
Per discussion, the new extensible node framework is thought to be
better designed than the custom path/scan/scanstate stuff we added
in PostgreSQL 9.5. Rework the latter to be more like the former.
This is not backward-compatible, but we generally don't promise that
for C APIs, and there probably aren't many people using this yet
anyway.
KaiGai Kohei, reviewed by Petr Jelinek and me. Some further
cosmetic changes by me.
Commit fbe5a3fb73 accidentally changed
this behavior; put things back the way they were, and add some
regression tests.
Report by Andres Freund; patch by Ashutosh Bapat, with a bit of
kibitzing by me.
cost_subplan() supposed that the given subplan must have plan_rows > 0,
which as far as I can tell was true until recent refactoring of the
code in createplan.c; but now that code allows the Result for a provably
empty subquery to have plan_rows = 0. Rather than undo that change,
put in a clamp to prevent zero divide.
get_cheapest_fractional_path() likewise supposed that best_path->rows > 0.
This assumption has been wrong for longer. It's actually harmless given
IEEE float math, because a positive value divided by zero gives +Infinity
and compare_fractional_path_costs() will do the right thing with that.
Still, best not to assume that.
final_cost_nestloop() also seems to have some risks in this area, so
borrow the clamping logic already present in the mergejoin cost functions.
Lastly, remove unnecessary clamp_row_est() in planner.c's calls to
get_number_of_groups(). The only thing that function does with path_rows
is pass it to estimate_num_groups() which already has an internal clamp,
so we don't need the extra call; and if we did, the callers are arguably
the wrong place for it anyway.
First two items reported by Piotr Stefaniak, the others are products
of my nosing around for similar problems. No back-patch since there's
no evidence that problems arise in the back branches.
In commit 9118d03a8c we taught the planner to postpone evaluation of
set-returning functions in a SELECT's targetlist until after any sort done
to satisfy ORDER BY. However, if we postpone some SRFs this way while
others do not get postponed (because they're sort or group key columns)
we will break the traditional behavior by which all SRFs in the tlist run
in-step during ExecTargetList(), so that you get the least common multiple
of their periods not the product. Fix make_sort_input_target() so it will
not split up SRF evaluation in such cases.
There is still a hazard of similar odd behavior if there's a SRF in a
grouping column and another one that isn't, but that was true before
and we're just trying to preserve bug-compatibility with the traditional
behavior. This whole area is overdue to be rethought and reimplemented,
but we'll try to avoid changing behavior until then.
Per report from Regina Obe.
Parallel workers can now partially aggregate the data and pass the
transition values back to the leader, which can combine the partial
results to produce the final answer.
David Rowley, based on earlier work by Haribabu Kommi. Reviewed by
Álvaro Herrera, Tomas Vondra, Amit Kapila, James Sewell, and me.
postgres_fdw can now sent an UPDATE or DELETE statement directly to
the foreign server in simple cases, rather than sending a SELECT FOR
UPDATE statement and then updating or deleting rows one-by-one.
Etsuro Fujita, reviewed by Rushabh Lathia, Shigeru Hanada, Kyotaro
Horiguchi, Albe Laurenz, Thom Brown, and me.
This means that, for example, "SELECT expensive_func(a) FROM bigtab
WHERE something" can compute expensive_func(a) in the workers rather
than the leader if it happens to be parallel-safe, which figures to be
a big win in some practical cases.
Currently, we can only do this if the entire target list is
parallel-safe. If we worked harder, we might be able to evaluate
parallel-safe targets in the worker and any parallel-restricted
targets in the leader, but that would be more complicated, and there
aren't that many parallel-restricted functions that people are likely
to use in queries anyway. I think. So just do the simple thing for
the moment.
Robert Haas, Amit Kapila, and Tom Lane
This is basically like the just-added create_upper_paths_hook, but
control is funneled only to the FDW responsible for all the baserels
of the current query; so providing such a callback is much less likely
to add useless overhead than using the hook function is.
The documentation is a bit sketchy. We'll likely want to improve it,
and/or adjust the call conventions, when we get some experience with
actually using this callback. Hopefully somebody will find time to
experiment with it before 9.6 feature freeze.
In the initial revision of the upper-planner pathification work, the only
available way for an FDW or custom-scan provider to inject Paths
representing post-scan-join processing was to insert them during scan-level
GetForeignPaths or similar processing. While that's not impossible, it'd
require quite a lot of duplicative processing to look forward and see if
the extension would be capable of implementing the whole query. To improve
matters for custom-scan providers, provide a hook function at the point
where the core code is about to start filling in upperrel Paths. At this
point Paths are available for the whole scan/join tree, which should reduce
the amount of redundant effort considerably.
(An alternative design that was suggested was to provide a separate hook
for each post-scan-join processing step, but that seems messy and not
clearly more useful.)
Following our time-honored tradition, there's no documentation for this
hook outside the source code.
As-is, this hook is only meant for custom scan providers, which we can't
assume very much about. A followon patch will implement an FDW callback
to let FDWs do the same thing in a somewhat more structured fashion.
Although the default choice of rel->reltarget should typically be
sufficient for scan or join paths, it's not at all sufficient for the
purposes PathTargets were invented for; in particular not for
upper-relation Paths. So break API compatibility by adding a PathTarget
argument to create_foreignscan_path(). To ease updating of existing
code, accept a NULL value of the argument as selecting rel->reltarget.
In commit 19a541143a I did not make PathTarget a subtype of Node,
and embedded a RelOptInfo's reltarget directly into it rather than having
a separately-allocated Node. In hindsight that was misguided
micro-optimization, enabled by the fact that at that point we didn't have
any Paths with custom PathTargets. Now that PathTarget processing has
been fleshed out some more, it's easier to see that it's better to have
PathTarget as an indepedent Node type, even if it does cost us one more
palloc to create a RelOptInfo. So change it while we still can.
This commit just changes the representation, without doing anything more
interesting than that.
CitusDB is using these and don't wish to redesign their code right now.
I am not on board with this being a good idea, or a good precedent,
but I lack the energy to fight about it.
It is frequently useful for volatile, set-returning, or expensive functions
in a SELECT's targetlist to be postponed till after ORDER BY and LIMIT are
done. Otherwise, the functions might be executed for every row of the
table despite the presence of LIMIT, and/or be executed in an unexpected
order. For example, in
SELECT x, nextval('seq') FROM tab ORDER BY x LIMIT 10;
it's probably desirable that the nextval() values are ordered the same
as x, and that nextval() is not run more than 10 times.
In the past, Postgres was inconsistent in this area: you would get the
desirable behavior if the ordering were performed via an indexscan, but
not if it had to be done by an explicit sort step. Getting the desired
behavior reliably required contortions like
SELECT x, nextval('seq')
FROM (SELECT x FROM tab ORDER BY x) ss LIMIT 10;
This patch conditionally postpones evaluation of pure-output target
expressions (that is, those that are not used as DISTINCT, ORDER BY, or
GROUP BY columns) so that they effectively occur after sorting, even if an
explicit sort step is necessary. Volatile expressions and set-returning
expressions are always postponed, so as to provide consistent semantics.
Expensive expressions (costing more than 10 times typical operator cost,
which by default would include any user-defined function) are postponed
if there is a LIMIT or if there are expressions that must be postponed.
We could be more aggressive and postpone any nontrivial expression, but
there are costs associated with doing so: it requires an extra Result plan
node which adds some overhead, and postponement changes the volume of data
going through the sort step, perhaps for the worse. Since we tend not to
have very good estimates of the output width of nontrivial expressions,
it's hard to have much confidence in our ability to predict whether
postponement would increase or decrease the cost of the sort; therefore
this patch doesn't attempt to make decisions conditionally on that.
Between these factors and a general desire not to change query behavior
when there's not a demonstrable benefit, it seems best to be conservative
about applying postponement. We might tweak the decision rules in the
future, though.
Konstantin Knizhnik, heavily rewritten by me
Teach make_group_input_target() and make_window_input_target() to work
entirely with the PathTarget representation of tlists, rather than
constructing a tlist and immediately deconstructing it into PathTarget
format. In itself this only saves a few palloc's; the bigger picture is
that it opens the door for sharing cost_qual_eval work across all of
planner.c's constructions of PathTargets. I'll come back to that later.
In support of this, flesh out tlist.c's infrastructure for PathTargets
a bit more.
All along, this function should have treated WindowFuncs in a manner
similar to Aggrefs, ie with an option whether or not to recurse into them.
By not considering the case, it was always recursing, which is OK for most
callers (although I suspect that the case in prepare_sort_from_pathkeys
might represent a bug). But now we need return-without-recursing behavior
as well. There are also more than a few callers that should never see a
WindowFunc, and now we'll get some error checking on that.
In commit 1d97c19a0f and later c1d9579dd8, we extended
pull_var_clause's API by adding enum-type arguments. That's sort of a pain
to maintain, though, because it means every time we add a new behavior we
must touch every last one of the call sites, even if there's a reasonable
default behavior that most of them could use. Let's switch over to using a
bitmask of flags, instead; that seems more maintainable and might save a
nanosecond or two as well. This commit changes no behavior in itself,
though I'm going to follow it up with one that does add a new behavior.
In passing, remove flatten_tlist(), which has not been used since 9.1
and would otherwise need the same API changes.
Removing these enums means that optimizer/tlist.h no longer needs to
depend on optimizer/var.h. Changing that caused a number of C files to
need addition of #include "optimizer/var.h" (probably we can thank old
runs of pgrminclude for that); but on balance it seems like a good change
anyway.
This function is written as though Gather doesn't project; but it does.
Even if it did not project, though, we must use build_path_tlist to ensure
that the output columns receive correct sortgroupref labeling.
Per report from Amit Kapila.
Refactor so that the internal APIs in planner.c deal in PathTargets not
targetlists, and establish a more regular structure for deriving the
targets needed for successive steps.
There is more that could be done here; calculating the eval costs of each
successive target independently is both inefficient and wrong in detail,
since we won't actually recompute values available from the input node's
tlist. But it's no worse than what happened before the pathification
rewrite. In any case this seems like a good starting point for considering
how to handle Konstantin Knizhnik's function-evaluation-postponement patch.
Instead of having planner.c compute a groupColIdx array and store it in
GroupingSetsPaths, make create_groupingsets_plan() find the grouping
columns by searching in the child plan node's tlist. Although that's
probably a bit slower for create_groupingsets_plan(), it's more like
the way every other plan node type does this, and it provides positive
confirmation that we know which child output columns we're supposed to be
grouping on. (Indeed, looking at this now, I'm not at all sure that it
wasn't broken before, because create_groupingsets_plan() isn't demanding
an exact tlist match from its child node.) Also, this allows substantial
simplification in planner.c, because it no longer needs to compute the
groupColIdx array at all; no other cases were using it.
I'd intended to put off this refactoring until later (like 9.7), but
in view of the likely bug fix and the need to rationalize planner.c's
tlist handling so we can do something sane with Konstantin Knizhnik's
function-evaluation-postponement patch, I think it can't wait.
I passed the wrong "root" struct to create_pathtarget in build_minmax_path.
Since the subroot is a clone of the outer root, this would not cause any
serious problems, but it would waste some cycles because
set_pathtarget_cost_width would not have access to Var width estimates
set up while running query_planner on the subroot.
This patch removes some redundant cost calculations that I left for later
cleanup in commit 3fc6e2d7f5. There's now a uniform policy that the
make_foo() convenience functions don't do any cost calculations. Most of
their callers copy costs from the source Path node, and for those that
don't, the calculation in the make_foo() function wasn't necessarily right
anyhow. (make_result() was particularly a mess, as it was serving multiple
callers using cost calcs designed for only the first one or two that had
ever existed.) Aside from saving a few cycles, this ensures that what
EXPLAIN prints matches the costs we used for planning purposes. It does
not change any planner decisions, since the decisions are already made.
Tom Lane
Robert Haas
Bruce Momjian
Alvaro Herrera
Simon Riggs
Andrew Gierth
Kevin Grittner
Peter Eisentraut
Andres Freund
Heikki Linnakangas
Greg Stark
Magnus Hagander
Teodor Sigaev