6257 lines
203 KiB
C
6257 lines
203 KiB
C
/*-------------------------------------------------------------------------
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*
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* costsize.c
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* Routines to compute (and set) relation sizes and path costs
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*
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* Path costs are measured in arbitrary units established by these basic
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* parameters:
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*
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* seq_page_cost Cost of a sequential page fetch
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* random_page_cost Cost of a non-sequential page fetch
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* cpu_tuple_cost Cost of typical CPU time to process a tuple
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* cpu_index_tuple_cost Cost of typical CPU time to process an index tuple
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* cpu_operator_cost Cost of CPU time to execute an operator or function
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* parallel_tuple_cost Cost of CPU time to pass a tuple from worker to leader backend
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* parallel_setup_cost Cost of setting up shared memory for parallelism
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*
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* We expect that the kernel will typically do some amount of read-ahead
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* optimization; this in conjunction with seek costs means that seq_page_cost
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* is normally considerably less than random_page_cost. (However, if the
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* database is fully cached in RAM, it is reasonable to set them equal.)
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*
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* We also use a rough estimate "effective_cache_size" of the number of
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* disk pages in Postgres + OS-level disk cache. (We can't simply use
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* NBuffers for this purpose because that would ignore the effects of
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* the kernel's disk cache.)
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*
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* Obviously, taking constants for these values is an oversimplification,
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* but it's tough enough to get any useful estimates even at this level of
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* detail. Note that all of these parameters are user-settable, in case
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* the default values are drastically off for a particular platform.
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*
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* seq_page_cost and random_page_cost can also be overridden for an individual
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* tablespace, in case some data is on a fast disk and other data is on a slow
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* disk. Per-tablespace overrides never apply to temporary work files such as
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* an external sort or a materialize node that overflows work_mem.
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*
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* We compute two separate costs for each path:
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* total_cost: total estimated cost to fetch all tuples
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* startup_cost: cost that is expended before first tuple is fetched
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* In some scenarios, such as when there is a LIMIT or we are implementing
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* an EXISTS(...) sub-select, it is not necessary to fetch all tuples of the
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* path's result. A caller can estimate the cost of fetching a partial
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* result by interpolating between startup_cost and total_cost. In detail:
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* actual_cost = startup_cost +
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* (total_cost - startup_cost) * tuples_to_fetch / path->rows;
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* Note that a base relation's rows count (and, by extension, plan_rows for
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* plan nodes below the LIMIT node) are set without regard to any LIMIT, so
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* that this equation works properly. (Note: while path->rows is never zero
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* for ordinary relations, it is zero for paths for provably-empty relations,
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* so beware of division-by-zero.) The LIMIT is applied as a top-level
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* plan node.
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*
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* For largely historical reasons, most of the routines in this module use
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* the passed result Path only to store their results (rows, startup_cost and
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* total_cost) into. All the input data they need is passed as separate
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* parameters, even though much of it could be extracted from the Path.
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* An exception is made for the cost_XXXjoin() routines, which expect all
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* the other fields of the passed XXXPath to be filled in, and similarly
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* cost_index() assumes the passed IndexPath is valid except for its output
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* values.
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*
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*
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* Portions Copyright (c) 1996-2023, PostgreSQL Global Development Group
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* Portions Copyright (c) 1994, Regents of the University of California
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*
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* IDENTIFICATION
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* src/backend/optimizer/path/costsize.c
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*
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*-------------------------------------------------------------------------
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*/
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#include "postgres.h"
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#include <limits.h>
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#include <math.h>
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#include "access/amapi.h"
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#include "access/htup_details.h"
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#include "access/tsmapi.h"
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#include "executor/executor.h"
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#include "executor/nodeAgg.h"
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#include "executor/nodeHash.h"
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#include "executor/nodeMemoize.h"
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#include "miscadmin.h"
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#include "nodes/makefuncs.h"
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#include "nodes/nodeFuncs.h"
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#include "optimizer/clauses.h"
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#include "optimizer/cost.h"
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#include "optimizer/optimizer.h"
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#include "optimizer/pathnode.h"
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#include "optimizer/paths.h"
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#include "optimizer/placeholder.h"
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#include "optimizer/plancat.h"
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#include "optimizer/planmain.h"
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#include "optimizer/restrictinfo.h"
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#include "parser/parsetree.h"
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#include "utils/lsyscache.h"
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#include "utils/selfuncs.h"
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#include "utils/spccache.h"
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#include "utils/tuplesort.h"
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#define LOG2(x) (log(x) / 0.693147180559945)
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/*
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* Append and MergeAppend nodes are less expensive than some other operations
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* which use cpu_tuple_cost; instead of adding a separate GUC, estimate the
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* per-tuple cost as cpu_tuple_cost multiplied by this value.
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*/
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#define APPEND_CPU_COST_MULTIPLIER 0.5
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/*
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* Maximum value for row estimates. We cap row estimates to this to help
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* ensure that costs based on these estimates remain within the range of what
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* double can represent. add_path() wouldn't act sanely given infinite or NaN
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* cost values.
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*/
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#define MAXIMUM_ROWCOUNT 1e100
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double seq_page_cost = DEFAULT_SEQ_PAGE_COST;
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double random_page_cost = DEFAULT_RANDOM_PAGE_COST;
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double cpu_tuple_cost = DEFAULT_CPU_TUPLE_COST;
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double cpu_index_tuple_cost = DEFAULT_CPU_INDEX_TUPLE_COST;
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double cpu_operator_cost = DEFAULT_CPU_OPERATOR_COST;
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double parallel_tuple_cost = DEFAULT_PARALLEL_TUPLE_COST;
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double parallel_setup_cost = DEFAULT_PARALLEL_SETUP_COST;
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double recursive_worktable_factor = DEFAULT_RECURSIVE_WORKTABLE_FACTOR;
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int effective_cache_size = DEFAULT_EFFECTIVE_CACHE_SIZE;
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Cost disable_cost = 1.0e10;
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int max_parallel_workers_per_gather = 2;
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bool enable_seqscan = true;
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bool enable_indexscan = true;
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bool enable_indexonlyscan = true;
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bool enable_bitmapscan = true;
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bool enable_tidscan = true;
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bool enable_sort = true;
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bool enable_incremental_sort = true;
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bool enable_hashagg = true;
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bool enable_nestloop = true;
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bool enable_material = true;
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bool enable_memoize = true;
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bool enable_mergejoin = true;
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bool enable_hashjoin = true;
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bool enable_gathermerge = true;
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bool enable_partitionwise_join = false;
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bool enable_partitionwise_aggregate = false;
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bool enable_parallel_append = true;
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bool enable_parallel_hash = true;
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bool enable_partition_pruning = true;
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bool enable_presorted_aggregate = true;
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bool enable_async_append = true;
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typedef struct
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{
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PlannerInfo *root;
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QualCost total;
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} cost_qual_eval_context;
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static List *extract_nonindex_conditions(List *qual_clauses, List *indexclauses);
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static MergeScanSelCache *cached_scansel(PlannerInfo *root,
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RestrictInfo *rinfo,
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PathKey *pathkey);
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static void cost_rescan(PlannerInfo *root, Path *path,
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Cost *rescan_startup_cost, Cost *rescan_total_cost);
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static bool cost_qual_eval_walker(Node *node, cost_qual_eval_context *context);
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static void get_restriction_qual_cost(PlannerInfo *root, RelOptInfo *baserel,
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ParamPathInfo *param_info,
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QualCost *qpqual_cost);
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static bool has_indexed_join_quals(NestPath *path);
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static double approx_tuple_count(PlannerInfo *root, JoinPath *path,
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List *quals);
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static double calc_joinrel_size_estimate(PlannerInfo *root,
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RelOptInfo *joinrel,
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RelOptInfo *outer_rel,
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RelOptInfo *inner_rel,
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double outer_rows,
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double inner_rows,
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SpecialJoinInfo *sjinfo,
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List *restrictlist);
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static Selectivity get_foreign_key_join_selectivity(PlannerInfo *root,
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Relids outer_relids,
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Relids inner_relids,
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SpecialJoinInfo *sjinfo,
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List **restrictlist);
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static Cost append_nonpartial_cost(List *subpaths, int numpaths,
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int parallel_workers);
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static void set_rel_width(PlannerInfo *root, RelOptInfo *rel);
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static int32 get_expr_width(PlannerInfo *root, const Node *expr);
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static double relation_byte_size(double tuples, int width);
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static double page_size(double tuples, int width);
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static double get_parallel_divisor(Path *path);
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/*
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* clamp_row_est
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* Force a row-count estimate to a sane value.
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*/
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double
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clamp_row_est(double nrows)
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{
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/*
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* Avoid infinite and NaN row estimates. Costs derived from such values
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* are going to be useless. Also force the estimate to be at least one
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* row, to make explain output look better and to avoid possible
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* divide-by-zero when interpolating costs. Make it an integer, too.
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*/
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if (nrows > MAXIMUM_ROWCOUNT || isnan(nrows))
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nrows = MAXIMUM_ROWCOUNT;
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else if (nrows <= 1.0)
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nrows = 1.0;
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else
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nrows = rint(nrows);
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return nrows;
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}
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/*
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* clamp_cardinality_to_long
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* Cast a Cardinality value to a sane long value.
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*/
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long
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clamp_cardinality_to_long(Cardinality x)
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{
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/*
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* Just for paranoia's sake, ensure we do something sane with negative or
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* NaN values.
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*/
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if (isnan(x))
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return LONG_MAX;
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if (x <= 0)
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return 0;
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/*
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* If "long" is 64 bits, then LONG_MAX cannot be represented exactly as a
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* double. Casting it to double and back may well result in overflow due
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* to rounding, so avoid doing that. We trust that any double value that
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* compares strictly less than "(double) LONG_MAX" will cast to a
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* representable "long" value.
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*/
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return (x < (double) LONG_MAX) ? (long) x : LONG_MAX;
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}
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/*
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* cost_seqscan
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* Determines and returns the cost of scanning a relation sequentially.
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*
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* 'baserel' is the relation to be scanned
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* 'param_info' is the ParamPathInfo if this is a parameterized path, else NULL
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*/
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void
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cost_seqscan(Path *path, PlannerInfo *root,
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RelOptInfo *baserel, ParamPathInfo *param_info)
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{
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Cost startup_cost = 0;
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Cost cpu_run_cost;
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Cost disk_run_cost;
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double spc_seq_page_cost;
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QualCost qpqual_cost;
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Cost cpu_per_tuple;
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/* Should only be applied to base relations */
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Assert(baserel->relid > 0);
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Assert(baserel->rtekind == RTE_RELATION);
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/* Mark the path with the correct row estimate */
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if (param_info)
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path->rows = param_info->ppi_rows;
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else
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path->rows = baserel->rows;
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if (!enable_seqscan)
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startup_cost += disable_cost;
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/* fetch estimated page cost for tablespace containing table */
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get_tablespace_page_costs(baserel->reltablespace,
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NULL,
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&spc_seq_page_cost);
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/*
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* disk costs
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*/
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disk_run_cost = spc_seq_page_cost * baserel->pages;
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/* CPU costs */
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get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
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startup_cost += qpqual_cost.startup;
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cpu_per_tuple = cpu_tuple_cost + qpqual_cost.per_tuple;
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cpu_run_cost = cpu_per_tuple * baserel->tuples;
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/* tlist eval costs are paid per output row, not per tuple scanned */
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startup_cost += path->pathtarget->cost.startup;
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cpu_run_cost += path->pathtarget->cost.per_tuple * path->rows;
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/* Adjust costing for parallelism, if used. */
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if (path->parallel_workers > 0)
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{
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double parallel_divisor = get_parallel_divisor(path);
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/* The CPU cost is divided among all the workers. */
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cpu_run_cost /= parallel_divisor;
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/*
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* It may be possible to amortize some of the I/O cost, but probably
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* not very much, because most operating systems already do aggressive
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* prefetching. For now, we assume that the disk run cost can't be
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* amortized at all.
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*/
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/*
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* In the case of a parallel plan, the row count needs to represent
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* the number of tuples processed per worker.
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*/
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path->rows = clamp_row_est(path->rows / parallel_divisor);
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}
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path->startup_cost = startup_cost;
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path->total_cost = startup_cost + cpu_run_cost + disk_run_cost;
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}
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/*
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* cost_samplescan
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* Determines and returns the cost of scanning a relation using sampling.
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*
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* 'baserel' is the relation to be scanned
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* 'param_info' is the ParamPathInfo if this is a parameterized path, else NULL
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*/
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void
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cost_samplescan(Path *path, PlannerInfo *root,
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RelOptInfo *baserel, ParamPathInfo *param_info)
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{
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Cost startup_cost = 0;
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Cost run_cost = 0;
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RangeTblEntry *rte;
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TableSampleClause *tsc;
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TsmRoutine *tsm;
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double spc_seq_page_cost,
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spc_random_page_cost,
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spc_page_cost;
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QualCost qpqual_cost;
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Cost cpu_per_tuple;
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/* Should only be applied to base relations with tablesample clauses */
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Assert(baserel->relid > 0);
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rte = planner_rt_fetch(baserel->relid, root);
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Assert(rte->rtekind == RTE_RELATION);
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tsc = rte->tablesample;
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Assert(tsc != NULL);
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tsm = GetTsmRoutine(tsc->tsmhandler);
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/* Mark the path with the correct row estimate */
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if (param_info)
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path->rows = param_info->ppi_rows;
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else
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path->rows = baserel->rows;
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/* fetch estimated page cost for tablespace containing table */
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get_tablespace_page_costs(baserel->reltablespace,
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&spc_random_page_cost,
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&spc_seq_page_cost);
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/* if NextSampleBlock is used, assume random access, else sequential */
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spc_page_cost = (tsm->NextSampleBlock != NULL) ?
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spc_random_page_cost : spc_seq_page_cost;
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/*
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* disk costs (recall that baserel->pages has already been set to the
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* number of pages the sampling method will visit)
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*/
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run_cost += spc_page_cost * baserel->pages;
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/*
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* CPU costs (recall that baserel->tuples has already been set to the
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* number of tuples the sampling method will select). Note that we ignore
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* execution cost of the TABLESAMPLE parameter expressions; they will be
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* evaluated only once per scan, and in most usages they'll likely be
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* simple constants anyway. We also don't charge anything for the
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* calculations the sampling method might do internally.
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*/
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get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
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startup_cost += qpqual_cost.startup;
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cpu_per_tuple = cpu_tuple_cost + qpqual_cost.per_tuple;
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run_cost += cpu_per_tuple * baserel->tuples;
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/* tlist eval costs are paid per output row, not per tuple scanned */
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startup_cost += path->pathtarget->cost.startup;
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run_cost += path->pathtarget->cost.per_tuple * path->rows;
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path->startup_cost = startup_cost;
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path->total_cost = startup_cost + run_cost;
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}
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/*
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* cost_gather
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* Determines and returns the cost of gather path.
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*
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* 'rel' is the relation to be operated upon
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* 'param_info' is the ParamPathInfo if this is a parameterized path, else NULL
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* 'rows' may be used to point to a row estimate; if non-NULL, it overrides
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* both 'rel' and 'param_info'. This is useful when the path doesn't exactly
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* correspond to any particular RelOptInfo.
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*/
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void
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cost_gather(GatherPath *path, PlannerInfo *root,
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RelOptInfo *rel, ParamPathInfo *param_info,
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double *rows)
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{
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Cost startup_cost = 0;
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Cost run_cost = 0;
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/* Mark the path with the correct row estimate */
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if (rows)
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path->path.rows = *rows;
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else if (param_info)
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path->path.rows = param_info->ppi_rows;
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else
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path->path.rows = rel->rows;
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startup_cost = path->subpath->startup_cost;
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run_cost = path->subpath->total_cost - path->subpath->startup_cost;
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/* Parallel setup and communication cost. */
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startup_cost += parallel_setup_cost;
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run_cost += parallel_tuple_cost * path->path.rows;
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path->path.startup_cost = startup_cost;
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path->path.total_cost = (startup_cost + run_cost);
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}
|
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|
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/*
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* cost_gather_merge
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* Determines and returns the cost of gather merge path.
|
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*
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* GatherMerge merges several pre-sorted input streams, using a heap that at
|
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* any given instant holds the next tuple from each stream. If there are N
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* streams, we need about N*log2(N) tuple comparisons to construct the heap at
|
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* startup, and then for each output tuple, about log2(N) comparisons to
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* replace the top heap entry with the next tuple from the same stream.
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*/
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void
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cost_gather_merge(GatherMergePath *path, PlannerInfo *root,
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RelOptInfo *rel, ParamPathInfo *param_info,
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Cost input_startup_cost, Cost input_total_cost,
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double *rows)
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{
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Cost startup_cost = 0;
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Cost run_cost = 0;
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Cost comparison_cost;
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double N;
|
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double logN;
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/* Mark the path with the correct row estimate */
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if (rows)
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path->path.rows = *rows;
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else if (param_info)
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path->path.rows = param_info->ppi_rows;
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else
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path->path.rows = rel->rows;
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|
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if (!enable_gathermerge)
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startup_cost += disable_cost;
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|
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/*
|
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* Add one to the number of workers to account for the leader. This might
|
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* be overgenerous since the leader will do less work than other workers
|
|
* in typical cases, but we'll go with it for now.
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*/
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Assert(path->num_workers > 0);
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N = (double) path->num_workers + 1;
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logN = LOG2(N);
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|
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/* Assumed cost per tuple comparison */
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comparison_cost = 2.0 * cpu_operator_cost;
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|
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/* Heap creation cost */
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startup_cost += comparison_cost * N * logN;
|
|
|
|
/* Per-tuple heap maintenance cost */
|
|
run_cost += path->path.rows * comparison_cost * logN;
|
|
|
|
/* small cost for heap management, like cost_merge_append */
|
|
run_cost += cpu_operator_cost * path->path.rows;
|
|
|
|
/*
|
|
* Parallel setup and communication cost. Since Gather Merge, unlike
|
|
* Gather, requires us to block until a tuple is available from every
|
|
* worker, we bump the IPC cost up a little bit as compared with Gather.
|
|
* For lack of a better idea, charge an extra 5%.
|
|
*/
|
|
startup_cost += parallel_setup_cost;
|
|
run_cost += parallel_tuple_cost * path->path.rows * 1.05;
|
|
|
|
path->path.startup_cost = startup_cost + input_startup_cost;
|
|
path->path.total_cost = (startup_cost + run_cost + input_total_cost);
|
|
}
|
|
|
|
/*
|
|
* cost_index
|
|
* Determines and returns the cost of scanning a relation using an index.
|
|
*
|
|
* 'path' describes the indexscan under consideration, and is complete
|
|
* except for the fields to be set by this routine
|
|
* 'loop_count' is the number of repetitions of the indexscan to factor into
|
|
* estimates of caching behavior
|
|
*
|
|
* In addition to rows, startup_cost and total_cost, cost_index() sets the
|
|
* path's indextotalcost and indexselectivity fields. These values will be
|
|
* needed if the IndexPath is used in a BitmapIndexScan.
|
|
*
|
|
* NOTE: path->indexquals must contain only clauses usable as index
|
|
* restrictions. Any additional quals evaluated as qpquals may reduce the
|
|
* number of returned tuples, but they won't reduce the number of tuples
|
|
* we have to fetch from the table, so they don't reduce the scan cost.
|
|
*/
|
|
void
|
|
cost_index(IndexPath *path, PlannerInfo *root, double loop_count,
|
|
bool partial_path)
|
|
{
|
|
IndexOptInfo *index = path->indexinfo;
|
|
RelOptInfo *baserel = index->rel;
|
|
bool indexonly = (path->path.pathtype == T_IndexOnlyScan);
|
|
amcostestimate_function amcostestimate;
|
|
List *qpquals;
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
Cost cpu_run_cost = 0;
|
|
Cost indexStartupCost;
|
|
Cost indexTotalCost;
|
|
Selectivity indexSelectivity;
|
|
double indexCorrelation,
|
|
csquared;
|
|
double spc_seq_page_cost,
|
|
spc_random_page_cost;
|
|
Cost min_IO_cost,
|
|
max_IO_cost;
|
|
QualCost qpqual_cost;
|
|
Cost cpu_per_tuple;
|
|
double tuples_fetched;
|
|
double pages_fetched;
|
|
double rand_heap_pages;
|
|
double index_pages;
|
|
|
|
/* Should only be applied to base relations */
|
|
Assert(IsA(baserel, RelOptInfo) &&
|
|
IsA(index, IndexOptInfo));
|
|
Assert(baserel->relid > 0);
|
|
Assert(baserel->rtekind == RTE_RELATION);
|
|
|
|
/*
|
|
* Mark the path with the correct row estimate, and identify which quals
|
|
* will need to be enforced as qpquals. We need not check any quals that
|
|
* are implied by the index's predicate, so we can use indrestrictinfo not
|
|
* baserestrictinfo as the list of relevant restriction clauses for the
|
|
* rel.
|
|
*/
|
|
if (path->path.param_info)
|
|
{
|
|
path->path.rows = path->path.param_info->ppi_rows;
|
|
/* qpquals come from the rel's restriction clauses and ppi_clauses */
|
|
qpquals = list_concat(extract_nonindex_conditions(path->indexinfo->indrestrictinfo,
|
|
path->indexclauses),
|
|
extract_nonindex_conditions(path->path.param_info->ppi_clauses,
|
|
path->indexclauses));
|
|
}
|
|
else
|
|
{
|
|
path->path.rows = baserel->rows;
|
|
/* qpquals come from just the rel's restriction clauses */
|
|
qpquals = extract_nonindex_conditions(path->indexinfo->indrestrictinfo,
|
|
path->indexclauses);
|
|
}
|
|
|
|
if (!enable_indexscan)
|
|
startup_cost += disable_cost;
|
|
/* we don't need to check enable_indexonlyscan; indxpath.c does that */
|
|
|
|
/*
|
|
* Call index-access-method-specific code to estimate the processing cost
|
|
* for scanning the index, as well as the selectivity of the index (ie,
|
|
* the fraction of main-table tuples we will have to retrieve) and its
|
|
* correlation to the main-table tuple order. We need a cast here because
|
|
* pathnodes.h uses a weak function type to avoid including amapi.h.
|
|
*/
|
|
amcostestimate = (amcostestimate_function) index->amcostestimate;
|
|
amcostestimate(root, path, loop_count,
|
|
&indexStartupCost, &indexTotalCost,
|
|
&indexSelectivity, &indexCorrelation,
|
|
&index_pages);
|
|
|
|
/*
|
|
* Save amcostestimate's results for possible use in bitmap scan planning.
|
|
* We don't bother to save indexStartupCost or indexCorrelation, because a
|
|
* bitmap scan doesn't care about either.
|
|
*/
|
|
path->indextotalcost = indexTotalCost;
|
|
path->indexselectivity = indexSelectivity;
|
|
|
|
/* all costs for touching index itself included here */
|
|
startup_cost += indexStartupCost;
|
|
run_cost += indexTotalCost - indexStartupCost;
|
|
|
|
/* estimate number of main-table tuples fetched */
|
|
tuples_fetched = clamp_row_est(indexSelectivity * baserel->tuples);
|
|
|
|
/* fetch estimated page costs for tablespace containing table */
|
|
get_tablespace_page_costs(baserel->reltablespace,
|
|
&spc_random_page_cost,
|
|
&spc_seq_page_cost);
|
|
|
|
/*----------
|
|
* Estimate number of main-table pages fetched, and compute I/O cost.
|
|
*
|
|
* When the index ordering is uncorrelated with the table ordering,
|
|
* we use an approximation proposed by Mackert and Lohman (see
|
|
* index_pages_fetched() for details) to compute the number of pages
|
|
* fetched, and then charge spc_random_page_cost per page fetched.
|
|
*
|
|
* When the index ordering is exactly correlated with the table ordering
|
|
* (just after a CLUSTER, for example), the number of pages fetched should
|
|
* be exactly selectivity * table_size. What's more, all but the first
|
|
* will be sequential fetches, not the random fetches that occur in the
|
|
* uncorrelated case. So if the number of pages is more than 1, we
|
|
* ought to charge
|
|
* spc_random_page_cost + (pages_fetched - 1) * spc_seq_page_cost
|
|
* For partially-correlated indexes, we ought to charge somewhere between
|
|
* these two estimates. We currently interpolate linearly between the
|
|
* estimates based on the correlation squared (XXX is that appropriate?).
|
|
*
|
|
* If it's an index-only scan, then we will not need to fetch any heap
|
|
* pages for which the visibility map shows all tuples are visible.
|
|
* Hence, reduce the estimated number of heap fetches accordingly.
|
|
* We use the measured fraction of the entire heap that is all-visible,
|
|
* which might not be particularly relevant to the subset of the heap
|
|
* that this query will fetch; but it's not clear how to do better.
|
|
*----------
|
|
*/
|
|
if (loop_count > 1)
|
|
{
|
|
/*
|
|
* For repeated indexscans, the appropriate estimate for the
|
|
* uncorrelated case is to scale up the number of tuples fetched in
|
|
* the Mackert and Lohman formula by the number of scans, so that we
|
|
* estimate the number of pages fetched by all the scans; then
|
|
* pro-rate the costs for one scan. In this case we assume all the
|
|
* fetches are random accesses.
|
|
*/
|
|
pages_fetched = index_pages_fetched(tuples_fetched * loop_count,
|
|
baserel->pages,
|
|
(double) index->pages,
|
|
root);
|
|
|
|
if (indexonly)
|
|
pages_fetched = ceil(pages_fetched * (1.0 - baserel->allvisfrac));
|
|
|
|
rand_heap_pages = pages_fetched;
|
|
|
|
max_IO_cost = (pages_fetched * spc_random_page_cost) / loop_count;
|
|
|
|
/*
|
|
* In the perfectly correlated case, the number of pages touched by
|
|
* each scan is selectivity * table_size, and we can use the Mackert
|
|
* and Lohman formula at the page level to estimate how much work is
|
|
* saved by caching across scans. We still assume all the fetches are
|
|
* random, though, which is an overestimate that's hard to correct for
|
|
* without double-counting the cache effects. (But in most cases
|
|
* where such a plan is actually interesting, only one page would get
|
|
* fetched per scan anyway, so it shouldn't matter much.)
|
|
*/
|
|
pages_fetched = ceil(indexSelectivity * (double) baserel->pages);
|
|
|
|
pages_fetched = index_pages_fetched(pages_fetched * loop_count,
|
|
baserel->pages,
|
|
(double) index->pages,
|
|
root);
|
|
|
|
if (indexonly)
|
|
pages_fetched = ceil(pages_fetched * (1.0 - baserel->allvisfrac));
|
|
|
|
min_IO_cost = (pages_fetched * spc_random_page_cost) / loop_count;
|
|
}
|
|
else
|
|
{
|
|
/*
|
|
* Normal case: apply the Mackert and Lohman formula, and then
|
|
* interpolate between that and the correlation-derived result.
|
|
*/
|
|
pages_fetched = index_pages_fetched(tuples_fetched,
|
|
baserel->pages,
|
|
(double) index->pages,
|
|
root);
|
|
|
|
if (indexonly)
|
|
pages_fetched = ceil(pages_fetched * (1.0 - baserel->allvisfrac));
|
|
|
|
rand_heap_pages = pages_fetched;
|
|
|
|
/* max_IO_cost is for the perfectly uncorrelated case (csquared=0) */
|
|
max_IO_cost = pages_fetched * spc_random_page_cost;
|
|
|
|
/* min_IO_cost is for the perfectly correlated case (csquared=1) */
|
|
pages_fetched = ceil(indexSelectivity * (double) baserel->pages);
|
|
|
|
if (indexonly)
|
|
pages_fetched = ceil(pages_fetched * (1.0 - baserel->allvisfrac));
|
|
|
|
if (pages_fetched > 0)
|
|
{
|
|
min_IO_cost = spc_random_page_cost;
|
|
if (pages_fetched > 1)
|
|
min_IO_cost += (pages_fetched - 1) * spc_seq_page_cost;
|
|
}
|
|
else
|
|
min_IO_cost = 0;
|
|
}
|
|
|
|
if (partial_path)
|
|
{
|
|
/*
|
|
* For index only scans compute workers based on number of index pages
|
|
* fetched; the number of heap pages we fetch might be so small as to
|
|
* effectively rule out parallelism, which we don't want to do.
|
|
*/
|
|
if (indexonly)
|
|
rand_heap_pages = -1;
|
|
|
|
/*
|
|
* Estimate the number of parallel workers required to scan index. Use
|
|
* the number of heap pages computed considering heap fetches won't be
|
|
* sequential as for parallel scans the pages are accessed in random
|
|
* order.
|
|
*/
|
|
path->path.parallel_workers = compute_parallel_worker(baserel,
|
|
rand_heap_pages,
|
|
index_pages,
|
|
max_parallel_workers_per_gather);
|
|
|
|
/*
|
|
* Fall out if workers can't be assigned for parallel scan, because in
|
|
* such a case this path will be rejected. So there is no benefit in
|
|
* doing extra computation.
|
|
*/
|
|
if (path->path.parallel_workers <= 0)
|
|
return;
|
|
|
|
path->path.parallel_aware = true;
|
|
}
|
|
|
|
/*
|
|
* Now interpolate based on estimated index order correlation to get total
|
|
* disk I/O cost for main table accesses.
|
|
*/
|
|
csquared = indexCorrelation * indexCorrelation;
|
|
|
|
run_cost += max_IO_cost + csquared * (min_IO_cost - max_IO_cost);
|
|
|
|
/*
|
|
* Estimate CPU costs per tuple.
|
|
*
|
|
* What we want here is cpu_tuple_cost plus the evaluation costs of any
|
|
* qual clauses that we have to evaluate as qpquals.
|
|
*/
|
|
cost_qual_eval(&qpqual_cost, qpquals, root);
|
|
|
|
startup_cost += qpqual_cost.startup;
|
|
cpu_per_tuple = cpu_tuple_cost + qpqual_cost.per_tuple;
|
|
|
|
cpu_run_cost += cpu_per_tuple * tuples_fetched;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->path.pathtarget->cost.startup;
|
|
cpu_run_cost += path->path.pathtarget->cost.per_tuple * path->path.rows;
|
|
|
|
/* Adjust costing for parallelism, if used. */
|
|
if (path->path.parallel_workers > 0)
|
|
{
|
|
double parallel_divisor = get_parallel_divisor(&path->path);
|
|
|
|
path->path.rows = clamp_row_est(path->path.rows / parallel_divisor);
|
|
|
|
/* The CPU cost is divided among all the workers. */
|
|
cpu_run_cost /= parallel_divisor;
|
|
}
|
|
|
|
run_cost += cpu_run_cost;
|
|
|
|
path->path.startup_cost = startup_cost;
|
|
path->path.total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* extract_nonindex_conditions
|
|
*
|
|
* Given a list of quals to be enforced in an indexscan, extract the ones that
|
|
* will have to be applied as qpquals (ie, the index machinery won't handle
|
|
* them). Here we detect only whether a qual clause is directly redundant
|
|
* with some indexclause. If the index path is chosen for use, createplan.c
|
|
* will try a bit harder to get rid of redundant qual conditions; specifically
|
|
* it will see if quals can be proven to be implied by the indexquals. But
|
|
* it does not seem worth the cycles to try to factor that in at this stage,
|
|
* since we're only trying to estimate qual eval costs. Otherwise this must
|
|
* match the logic in create_indexscan_plan().
|
|
*
|
|
* qual_clauses, and the result, are lists of RestrictInfos.
|
|
* indexclauses is a list of IndexClauses.
|
|
*/
|
|
static List *
|
|
extract_nonindex_conditions(List *qual_clauses, List *indexclauses)
|
|
{
|
|
List *result = NIL;
|
|
ListCell *lc;
|
|
|
|
foreach(lc, qual_clauses)
|
|
{
|
|
RestrictInfo *rinfo = lfirst_node(RestrictInfo, lc);
|
|
|
|
if (rinfo->pseudoconstant)
|
|
continue; /* we may drop pseudoconstants here */
|
|
if (is_redundant_with_indexclauses(rinfo, indexclauses))
|
|
continue; /* dup or derived from same EquivalenceClass */
|
|
/* ... skip the predicate proof attempt createplan.c will try ... */
|
|
result = lappend(result, rinfo);
|
|
}
|
|
return result;
|
|
}
|
|
|
|
/*
|
|
* index_pages_fetched
|
|
* Estimate the number of pages actually fetched after accounting for
|
|
* cache effects.
|
|
*
|
|
* We use an approximation proposed by Mackert and Lohman, "Index Scans
|
|
* Using a Finite LRU Buffer: A Validated I/O Model", ACM Transactions
|
|
* on Database Systems, Vol. 14, No. 3, September 1989, Pages 401-424.
|
|
* The Mackert and Lohman approximation is that the number of pages
|
|
* fetched is
|
|
* PF =
|
|
* min(2TNs/(2T+Ns), T) when T <= b
|
|
* 2TNs/(2T+Ns) when T > b and Ns <= 2Tb/(2T-b)
|
|
* b + (Ns - 2Tb/(2T-b))*(T-b)/T when T > b and Ns > 2Tb/(2T-b)
|
|
* where
|
|
* T = # pages in table
|
|
* N = # tuples in table
|
|
* s = selectivity = fraction of table to be scanned
|
|
* b = # buffer pages available (we include kernel space here)
|
|
*
|
|
* We assume that effective_cache_size is the total number of buffer pages
|
|
* available for the whole query, and pro-rate that space across all the
|
|
* tables in the query and the index currently under consideration. (This
|
|
* ignores space needed for other indexes used by the query, but since we
|
|
* don't know which indexes will get used, we can't estimate that very well;
|
|
* and in any case counting all the tables may well be an overestimate, since
|
|
* depending on the join plan not all the tables may be scanned concurrently.)
|
|
*
|
|
* The product Ns is the number of tuples fetched; we pass in that
|
|
* product rather than calculating it here. "pages" is the number of pages
|
|
* in the object under consideration (either an index or a table).
|
|
* "index_pages" is the amount to add to the total table space, which was
|
|
* computed for us by make_one_rel.
|
|
*
|
|
* Caller is expected to have ensured that tuples_fetched is greater than zero
|
|
* and rounded to integer (see clamp_row_est). The result will likewise be
|
|
* greater than zero and integral.
|
|
*/
|
|
double
|
|
index_pages_fetched(double tuples_fetched, BlockNumber pages,
|
|
double index_pages, PlannerInfo *root)
|
|
{
|
|
double pages_fetched;
|
|
double total_pages;
|
|
double T,
|
|
b;
|
|
|
|
/* T is # pages in table, but don't allow it to be zero */
|
|
T = (pages > 1) ? (double) pages : 1.0;
|
|
|
|
/* Compute number of pages assumed to be competing for cache space */
|
|
total_pages = root->total_table_pages + index_pages;
|
|
total_pages = Max(total_pages, 1.0);
|
|
Assert(T <= total_pages);
|
|
|
|
/* b is pro-rated share of effective_cache_size */
|
|
b = (double) effective_cache_size * T / total_pages;
|
|
|
|
/* force it positive and integral */
|
|
if (b <= 1.0)
|
|
b = 1.0;
|
|
else
|
|
b = ceil(b);
|
|
|
|
/* This part is the Mackert and Lohman formula */
|
|
if (T <= b)
|
|
{
|
|
pages_fetched =
|
|
(2.0 * T * tuples_fetched) / (2.0 * T + tuples_fetched);
|
|
if (pages_fetched >= T)
|
|
pages_fetched = T;
|
|
else
|
|
pages_fetched = ceil(pages_fetched);
|
|
}
|
|
else
|
|
{
|
|
double lim;
|
|
|
|
lim = (2.0 * T * b) / (2.0 * T - b);
|
|
if (tuples_fetched <= lim)
|
|
{
|
|
pages_fetched =
|
|
(2.0 * T * tuples_fetched) / (2.0 * T + tuples_fetched);
|
|
}
|
|
else
|
|
{
|
|
pages_fetched =
|
|
b + (tuples_fetched - lim) * (T - b) / T;
|
|
}
|
|
pages_fetched = ceil(pages_fetched);
|
|
}
|
|
return pages_fetched;
|
|
}
|
|
|
|
/*
|
|
* get_indexpath_pages
|
|
* Determine the total size of the indexes used in a bitmap index path.
|
|
*
|
|
* Note: if the same index is used more than once in a bitmap tree, we will
|
|
* count it multiple times, which perhaps is the wrong thing ... but it's
|
|
* not completely clear, and detecting duplicates is difficult, so ignore it
|
|
* for now.
|
|
*/
|
|
static double
|
|
get_indexpath_pages(Path *bitmapqual)
|
|
{
|
|
double result = 0;
|
|
ListCell *l;
|
|
|
|
if (IsA(bitmapqual, BitmapAndPath))
|
|
{
|
|
BitmapAndPath *apath = (BitmapAndPath *) bitmapqual;
|
|
|
|
foreach(l, apath->bitmapquals)
|
|
{
|
|
result += get_indexpath_pages((Path *) lfirst(l));
|
|
}
|
|
}
|
|
else if (IsA(bitmapqual, BitmapOrPath))
|
|
{
|
|
BitmapOrPath *opath = (BitmapOrPath *) bitmapqual;
|
|
|
|
foreach(l, opath->bitmapquals)
|
|
{
|
|
result += get_indexpath_pages((Path *) lfirst(l));
|
|
}
|
|
}
|
|
else if (IsA(bitmapqual, IndexPath))
|
|
{
|
|
IndexPath *ipath = (IndexPath *) bitmapqual;
|
|
|
|
result = (double) ipath->indexinfo->pages;
|
|
}
|
|
else
|
|
elog(ERROR, "unrecognized node type: %d", nodeTag(bitmapqual));
|
|
|
|
return result;
|
|
}
|
|
|
|
/*
|
|
* cost_bitmap_heap_scan
|
|
* Determines and returns the cost of scanning a relation using a bitmap
|
|
* index-then-heap plan.
|
|
*
|
|
* 'baserel' is the relation to be scanned
|
|
* 'param_info' is the ParamPathInfo if this is a parameterized path, else NULL
|
|
* 'bitmapqual' is a tree of IndexPaths, BitmapAndPaths, and BitmapOrPaths
|
|
* 'loop_count' is the number of repetitions of the indexscan to factor into
|
|
* estimates of caching behavior
|
|
*
|
|
* Note: the component IndexPaths in bitmapqual should have been costed
|
|
* using the same loop_count.
|
|
*/
|
|
void
|
|
cost_bitmap_heap_scan(Path *path, PlannerInfo *root, RelOptInfo *baserel,
|
|
ParamPathInfo *param_info,
|
|
Path *bitmapqual, double loop_count)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
Cost indexTotalCost;
|
|
QualCost qpqual_cost;
|
|
Cost cpu_per_tuple;
|
|
Cost cost_per_page;
|
|
Cost cpu_run_cost;
|
|
double tuples_fetched;
|
|
double pages_fetched;
|
|
double spc_seq_page_cost,
|
|
spc_random_page_cost;
|
|
double T;
|
|
|
|
/* Should only be applied to base relations */
|
|
Assert(IsA(baserel, RelOptInfo));
|
|
Assert(baserel->relid > 0);
|
|
Assert(baserel->rtekind == RTE_RELATION);
|
|
|
|
/* Mark the path with the correct row estimate */
|
|
if (param_info)
|
|
path->rows = param_info->ppi_rows;
|
|
else
|
|
path->rows = baserel->rows;
|
|
|
|
if (!enable_bitmapscan)
|
|
startup_cost += disable_cost;
|
|
|
|
pages_fetched = compute_bitmap_pages(root, baserel, bitmapqual,
|
|
loop_count, &indexTotalCost,
|
|
&tuples_fetched);
|
|
|
|
startup_cost += indexTotalCost;
|
|
T = (baserel->pages > 1) ? (double) baserel->pages : 1.0;
|
|
|
|
/* Fetch estimated page costs for tablespace containing table. */
|
|
get_tablespace_page_costs(baserel->reltablespace,
|
|
&spc_random_page_cost,
|
|
&spc_seq_page_cost);
|
|
|
|
/*
|
|
* For small numbers of pages we should charge spc_random_page_cost
|
|
* apiece, while if nearly all the table's pages are being read, it's more
|
|
* appropriate to charge spc_seq_page_cost apiece. The effect is
|
|
* nonlinear, too. For lack of a better idea, interpolate like this to
|
|
* determine the cost per page.
|
|
*/
|
|
if (pages_fetched >= 2.0)
|
|
cost_per_page = spc_random_page_cost -
|
|
(spc_random_page_cost - spc_seq_page_cost)
|
|
* sqrt(pages_fetched / T);
|
|
else
|
|
cost_per_page = spc_random_page_cost;
|
|
|
|
run_cost += pages_fetched * cost_per_page;
|
|
|
|
/*
|
|
* Estimate CPU costs per tuple.
|
|
*
|
|
* Often the indexquals don't need to be rechecked at each tuple ... but
|
|
* not always, especially not if there are enough tuples involved that the
|
|
* bitmaps become lossy. For the moment, just assume they will be
|
|
* rechecked always. This means we charge the full freight for all the
|
|
* scan clauses.
|
|
*/
|
|
get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
|
|
|
|
startup_cost += qpqual_cost.startup;
|
|
cpu_per_tuple = cpu_tuple_cost + qpqual_cost.per_tuple;
|
|
cpu_run_cost = cpu_per_tuple * tuples_fetched;
|
|
|
|
/* Adjust costing for parallelism, if used. */
|
|
if (path->parallel_workers > 0)
|
|
{
|
|
double parallel_divisor = get_parallel_divisor(path);
|
|
|
|
/* The CPU cost is divided among all the workers. */
|
|
cpu_run_cost /= parallel_divisor;
|
|
|
|
path->rows = clamp_row_est(path->rows / parallel_divisor);
|
|
}
|
|
|
|
|
|
run_cost += cpu_run_cost;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->pathtarget->cost.startup;
|
|
run_cost += path->pathtarget->cost.per_tuple * path->rows;
|
|
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_bitmap_tree_node
|
|
* Extract cost and selectivity from a bitmap tree node (index/and/or)
|
|
*/
|
|
void
|
|
cost_bitmap_tree_node(Path *path, Cost *cost, Selectivity *selec)
|
|
{
|
|
if (IsA(path, IndexPath))
|
|
{
|
|
*cost = ((IndexPath *) path)->indextotalcost;
|
|
*selec = ((IndexPath *) path)->indexselectivity;
|
|
|
|
/*
|
|
* Charge a small amount per retrieved tuple to reflect the costs of
|
|
* manipulating the bitmap. This is mostly to make sure that a bitmap
|
|
* scan doesn't look to be the same cost as an indexscan to retrieve a
|
|
* single tuple.
|
|
*/
|
|
*cost += 0.1 * cpu_operator_cost * path->rows;
|
|
}
|
|
else if (IsA(path, BitmapAndPath))
|
|
{
|
|
*cost = path->total_cost;
|
|
*selec = ((BitmapAndPath *) path)->bitmapselectivity;
|
|
}
|
|
else if (IsA(path, BitmapOrPath))
|
|
{
|
|
*cost = path->total_cost;
|
|
*selec = ((BitmapOrPath *) path)->bitmapselectivity;
|
|
}
|
|
else
|
|
{
|
|
elog(ERROR, "unrecognized node type: %d", nodeTag(path));
|
|
*cost = *selec = 0; /* keep compiler quiet */
|
|
}
|
|
}
|
|
|
|
/*
|
|
* cost_bitmap_and_node
|
|
* Estimate the cost of a BitmapAnd node
|
|
*
|
|
* Note that this considers only the costs of index scanning and bitmap
|
|
* creation, not the eventual heap access. In that sense the object isn't
|
|
* truly a Path, but it has enough path-like properties (costs in particular)
|
|
* to warrant treating it as one. We don't bother to set the path rows field,
|
|
* however.
|
|
*/
|
|
void
|
|
cost_bitmap_and_node(BitmapAndPath *path, PlannerInfo *root)
|
|
{
|
|
Cost totalCost;
|
|
Selectivity selec;
|
|
ListCell *l;
|
|
|
|
/*
|
|
* We estimate AND selectivity on the assumption that the inputs are
|
|
* independent. This is probably often wrong, but we don't have the info
|
|
* to do better.
|
|
*
|
|
* The runtime cost of the BitmapAnd itself is estimated at 100x
|
|
* cpu_operator_cost for each tbm_intersect needed. Probably too small,
|
|
* definitely too simplistic?
|
|
*/
|
|
totalCost = 0.0;
|
|
selec = 1.0;
|
|
foreach(l, path->bitmapquals)
|
|
{
|
|
Path *subpath = (Path *) lfirst(l);
|
|
Cost subCost;
|
|
Selectivity subselec;
|
|
|
|
cost_bitmap_tree_node(subpath, &subCost, &subselec);
|
|
|
|
selec *= subselec;
|
|
|
|
totalCost += subCost;
|
|
if (l != list_head(path->bitmapquals))
|
|
totalCost += 100.0 * cpu_operator_cost;
|
|
}
|
|
path->bitmapselectivity = selec;
|
|
path->path.rows = 0; /* per above, not used */
|
|
path->path.startup_cost = totalCost;
|
|
path->path.total_cost = totalCost;
|
|
}
|
|
|
|
/*
|
|
* cost_bitmap_or_node
|
|
* Estimate the cost of a BitmapOr node
|
|
*
|
|
* See comments for cost_bitmap_and_node.
|
|
*/
|
|
void
|
|
cost_bitmap_or_node(BitmapOrPath *path, PlannerInfo *root)
|
|
{
|
|
Cost totalCost;
|
|
Selectivity selec;
|
|
ListCell *l;
|
|
|
|
/*
|
|
* We estimate OR selectivity on the assumption that the inputs are
|
|
* non-overlapping, since that's often the case in "x IN (list)" type
|
|
* situations. Of course, we clamp to 1.0 at the end.
|
|
*
|
|
* The runtime cost of the BitmapOr itself is estimated at 100x
|
|
* cpu_operator_cost for each tbm_union needed. Probably too small,
|
|
* definitely too simplistic? We are aware that the tbm_unions are
|
|
* optimized out when the inputs are BitmapIndexScans.
|
|
*/
|
|
totalCost = 0.0;
|
|
selec = 0.0;
|
|
foreach(l, path->bitmapquals)
|
|
{
|
|
Path *subpath = (Path *) lfirst(l);
|
|
Cost subCost;
|
|
Selectivity subselec;
|
|
|
|
cost_bitmap_tree_node(subpath, &subCost, &subselec);
|
|
|
|
selec += subselec;
|
|
|
|
totalCost += subCost;
|
|
if (l != list_head(path->bitmapquals) &&
|
|
!IsA(subpath, IndexPath))
|
|
totalCost += 100.0 * cpu_operator_cost;
|
|
}
|
|
path->bitmapselectivity = Min(selec, 1.0);
|
|
path->path.rows = 0; /* per above, not used */
|
|
path->path.startup_cost = totalCost;
|
|
path->path.total_cost = totalCost;
|
|
}
|
|
|
|
/*
|
|
* cost_tidscan
|
|
* Determines and returns the cost of scanning a relation using TIDs.
|
|
*
|
|
* 'baserel' is the relation to be scanned
|
|
* 'tidquals' is the list of TID-checkable quals
|
|
* 'param_info' is the ParamPathInfo if this is a parameterized path, else NULL
|
|
*/
|
|
void
|
|
cost_tidscan(Path *path, PlannerInfo *root,
|
|
RelOptInfo *baserel, List *tidquals, ParamPathInfo *param_info)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
bool isCurrentOf = false;
|
|
QualCost qpqual_cost;
|
|
Cost cpu_per_tuple;
|
|
QualCost tid_qual_cost;
|
|
int ntuples;
|
|
ListCell *l;
|
|
double spc_random_page_cost;
|
|
|
|
/* Should only be applied to base relations */
|
|
Assert(baserel->relid > 0);
|
|
Assert(baserel->rtekind == RTE_RELATION);
|
|
|
|
/* Mark the path with the correct row estimate */
|
|
if (param_info)
|
|
path->rows = param_info->ppi_rows;
|
|
else
|
|
path->rows = baserel->rows;
|
|
|
|
/* Count how many tuples we expect to retrieve */
|
|
ntuples = 0;
|
|
foreach(l, tidquals)
|
|
{
|
|
RestrictInfo *rinfo = lfirst_node(RestrictInfo, l);
|
|
Expr *qual = rinfo->clause;
|
|
|
|
if (IsA(qual, ScalarArrayOpExpr))
|
|
{
|
|
/* Each element of the array yields 1 tuple */
|
|
ScalarArrayOpExpr *saop = (ScalarArrayOpExpr *) qual;
|
|
Node *arraynode = (Node *) lsecond(saop->args);
|
|
|
|
ntuples += estimate_array_length(arraynode);
|
|
}
|
|
else if (IsA(qual, CurrentOfExpr))
|
|
{
|
|
/* CURRENT OF yields 1 tuple */
|
|
isCurrentOf = true;
|
|
ntuples++;
|
|
}
|
|
else
|
|
{
|
|
/* It's just CTID = something, count 1 tuple */
|
|
ntuples++;
|
|
}
|
|
}
|
|
|
|
/*
|
|
* We must force TID scan for WHERE CURRENT OF, because only nodeTidscan.c
|
|
* understands how to do it correctly. Therefore, honor enable_tidscan
|
|
* only when CURRENT OF isn't present. Also note that cost_qual_eval
|
|
* counts a CurrentOfExpr as having startup cost disable_cost, which we
|
|
* subtract off here; that's to prevent other plan types such as seqscan
|
|
* from winning.
|
|
*/
|
|
if (isCurrentOf)
|
|
{
|
|
Assert(baserel->baserestrictcost.startup >= disable_cost);
|
|
startup_cost -= disable_cost;
|
|
}
|
|
else if (!enable_tidscan)
|
|
startup_cost += disable_cost;
|
|
|
|
/*
|
|
* The TID qual expressions will be computed once, any other baserestrict
|
|
* quals once per retrieved tuple.
|
|
*/
|
|
cost_qual_eval(&tid_qual_cost, tidquals, root);
|
|
|
|
/* fetch estimated page cost for tablespace containing table */
|
|
get_tablespace_page_costs(baserel->reltablespace,
|
|
&spc_random_page_cost,
|
|
NULL);
|
|
|
|
/* disk costs --- assume each tuple on a different page */
|
|
run_cost += spc_random_page_cost * ntuples;
|
|
|
|
/* Add scanning CPU costs */
|
|
get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
|
|
|
|
/* XXX currently we assume TID quals are a subset of qpquals */
|
|
startup_cost += qpqual_cost.startup + tid_qual_cost.per_tuple;
|
|
cpu_per_tuple = cpu_tuple_cost + qpqual_cost.per_tuple -
|
|
tid_qual_cost.per_tuple;
|
|
run_cost += cpu_per_tuple * ntuples;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->pathtarget->cost.startup;
|
|
run_cost += path->pathtarget->cost.per_tuple * path->rows;
|
|
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_tidrangescan
|
|
* Determines and sets the costs of scanning a relation using a range of
|
|
* TIDs for 'path'
|
|
*
|
|
* 'baserel' is the relation to be scanned
|
|
* 'tidrangequals' is the list of TID-checkable range quals
|
|
* 'param_info' is the ParamPathInfo if this is a parameterized path, else NULL
|
|
*/
|
|
void
|
|
cost_tidrangescan(Path *path, PlannerInfo *root,
|
|
RelOptInfo *baserel, List *tidrangequals,
|
|
ParamPathInfo *param_info)
|
|
{
|
|
Selectivity selectivity;
|
|
double pages;
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
QualCost qpqual_cost;
|
|
Cost cpu_per_tuple;
|
|
QualCost tid_qual_cost;
|
|
double ntuples;
|
|
double nseqpages;
|
|
double spc_random_page_cost;
|
|
double spc_seq_page_cost;
|
|
|
|
/* Should only be applied to base relations */
|
|
Assert(baserel->relid > 0);
|
|
Assert(baserel->rtekind == RTE_RELATION);
|
|
|
|
/* Mark the path with the correct row estimate */
|
|
if (param_info)
|
|
path->rows = param_info->ppi_rows;
|
|
else
|
|
path->rows = baserel->rows;
|
|
|
|
/* Count how many tuples and pages we expect to scan */
|
|
selectivity = clauselist_selectivity(root, tidrangequals, baserel->relid,
|
|
JOIN_INNER, NULL);
|
|
pages = ceil(selectivity * baserel->pages);
|
|
|
|
if (pages <= 0.0)
|
|
pages = 1.0;
|
|
|
|
/*
|
|
* The first page in a range requires a random seek, but each subsequent
|
|
* page is just a normal sequential page read. NOTE: it's desirable for
|
|
* TID Range Scans to cost more than the equivalent Sequential Scans,
|
|
* because Seq Scans have some performance advantages such as scan
|
|
* synchronization and parallelizability, and we'd prefer one of them to
|
|
* be picked unless a TID Range Scan really is better.
|
|
*/
|
|
ntuples = selectivity * baserel->tuples;
|
|
nseqpages = pages - 1.0;
|
|
|
|
if (!enable_tidscan)
|
|
startup_cost += disable_cost;
|
|
|
|
/*
|
|
* The TID qual expressions will be computed once, any other baserestrict
|
|
* quals once per retrieved tuple.
|
|
*/
|
|
cost_qual_eval(&tid_qual_cost, tidrangequals, root);
|
|
|
|
/* fetch estimated page cost for tablespace containing table */
|
|
get_tablespace_page_costs(baserel->reltablespace,
|
|
&spc_random_page_cost,
|
|
&spc_seq_page_cost);
|
|
|
|
/* disk costs; 1 random page and the remainder as seq pages */
|
|
run_cost += spc_random_page_cost + spc_seq_page_cost * nseqpages;
|
|
|
|
/* Add scanning CPU costs */
|
|
get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
|
|
|
|
/*
|
|
* XXX currently we assume TID quals are a subset of qpquals at this
|
|
* point; they will be removed (if possible) when we create the plan, so
|
|
* we subtract their cost from the total qpqual cost. (If the TID quals
|
|
* can't be removed, this is a mistake and we're going to underestimate
|
|
* the CPU cost a bit.)
|
|
*/
|
|
startup_cost += qpqual_cost.startup + tid_qual_cost.per_tuple;
|
|
cpu_per_tuple = cpu_tuple_cost + qpqual_cost.per_tuple -
|
|
tid_qual_cost.per_tuple;
|
|
run_cost += cpu_per_tuple * ntuples;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->pathtarget->cost.startup;
|
|
run_cost += path->pathtarget->cost.per_tuple * path->rows;
|
|
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_subqueryscan
|
|
* Determines and returns the cost of scanning a subquery RTE.
|
|
*
|
|
* 'baserel' is the relation to be scanned
|
|
* 'param_info' is the ParamPathInfo if this is a parameterized path, else NULL
|
|
* 'trivial_pathtarget' is true if the pathtarget is believed to be trivial.
|
|
*/
|
|
void
|
|
cost_subqueryscan(SubqueryScanPath *path, PlannerInfo *root,
|
|
RelOptInfo *baserel, ParamPathInfo *param_info,
|
|
bool trivial_pathtarget)
|
|
{
|
|
Cost startup_cost;
|
|
Cost run_cost;
|
|
List *qpquals;
|
|
QualCost qpqual_cost;
|
|
Cost cpu_per_tuple;
|
|
|
|
/* Should only be applied to base relations that are subqueries */
|
|
Assert(baserel->relid > 0);
|
|
Assert(baserel->rtekind == RTE_SUBQUERY);
|
|
|
|
/*
|
|
* We compute the rowcount estimate as the subplan's estimate times the
|
|
* selectivity of relevant restriction clauses. In simple cases this will
|
|
* come out the same as baserel->rows; but when dealing with parallelized
|
|
* paths we must do it like this to get the right answer.
|
|
*/
|
|
if (param_info)
|
|
qpquals = list_concat_copy(param_info->ppi_clauses,
|
|
baserel->baserestrictinfo);
|
|
else
|
|
qpquals = baserel->baserestrictinfo;
|
|
|
|
path->path.rows = clamp_row_est(path->subpath->rows *
|
|
clauselist_selectivity(root,
|
|
qpquals,
|
|
0,
|
|
JOIN_INNER,
|
|
NULL));
|
|
|
|
/*
|
|
* Cost of path is cost of evaluating the subplan, plus cost of evaluating
|
|
* any restriction clauses and tlist that will be attached to the
|
|
* SubqueryScan node, plus cpu_tuple_cost to account for selection and
|
|
* projection overhead.
|
|
*/
|
|
path->path.startup_cost = path->subpath->startup_cost;
|
|
path->path.total_cost = path->subpath->total_cost;
|
|
|
|
/*
|
|
* However, if there are no relevant restriction clauses and the
|
|
* pathtarget is trivial, then we expect that setrefs.c will optimize away
|
|
* the SubqueryScan plan node altogether, so we should just make its cost
|
|
* and rowcount equal to the input path's.
|
|
*
|
|
* Note: there are some edge cases where createplan.c will apply a
|
|
* different targetlist to the SubqueryScan node, thus falsifying our
|
|
* current estimate of whether the target is trivial, and making the cost
|
|
* estimate (though not the rowcount) wrong. It does not seem worth the
|
|
* extra complication to try to account for that exactly, especially since
|
|
* that behavior falsifies other cost estimates as well.
|
|
*/
|
|
if (qpquals == NIL && trivial_pathtarget)
|
|
return;
|
|
|
|
get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
|
|
|
|
startup_cost = qpqual_cost.startup;
|
|
cpu_per_tuple = cpu_tuple_cost + qpqual_cost.per_tuple;
|
|
run_cost = cpu_per_tuple * path->subpath->rows;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->path.pathtarget->cost.startup;
|
|
run_cost += path->path.pathtarget->cost.per_tuple * path->path.rows;
|
|
|
|
path->path.startup_cost += startup_cost;
|
|
path->path.total_cost += startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_functionscan
|
|
* Determines and returns the cost of scanning a function RTE.
|
|
*
|
|
* 'baserel' is the relation to be scanned
|
|
* 'param_info' is the ParamPathInfo if this is a parameterized path, else NULL
|
|
*/
|
|
void
|
|
cost_functionscan(Path *path, PlannerInfo *root,
|
|
RelOptInfo *baserel, ParamPathInfo *param_info)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
QualCost qpqual_cost;
|
|
Cost cpu_per_tuple;
|
|
RangeTblEntry *rte;
|
|
QualCost exprcost;
|
|
|
|
/* Should only be applied to base relations that are functions */
|
|
Assert(baserel->relid > 0);
|
|
rte = planner_rt_fetch(baserel->relid, root);
|
|
Assert(rte->rtekind == RTE_FUNCTION);
|
|
|
|
/* Mark the path with the correct row estimate */
|
|
if (param_info)
|
|
path->rows = param_info->ppi_rows;
|
|
else
|
|
path->rows = baserel->rows;
|
|
|
|
/*
|
|
* Estimate costs of executing the function expression(s).
|
|
*
|
|
* Currently, nodeFunctionscan.c always executes the functions to
|
|
* completion before returning any rows, and caches the results in a
|
|
* tuplestore. So the function eval cost is all startup cost, and per-row
|
|
* costs are minimal.
|
|
*
|
|
* XXX in principle we ought to charge tuplestore spill costs if the
|
|
* number of rows is large. However, given how phony our rowcount
|
|
* estimates for functions tend to be, there's not a lot of point in that
|
|
* refinement right now.
|
|
*/
|
|
cost_qual_eval_node(&exprcost, (Node *) rte->functions, root);
|
|
|
|
startup_cost += exprcost.startup + exprcost.per_tuple;
|
|
|
|
/* Add scanning CPU costs */
|
|
get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
|
|
|
|
startup_cost += qpqual_cost.startup;
|
|
cpu_per_tuple = cpu_tuple_cost + qpqual_cost.per_tuple;
|
|
run_cost += cpu_per_tuple * baserel->tuples;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->pathtarget->cost.startup;
|
|
run_cost += path->pathtarget->cost.per_tuple * path->rows;
|
|
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_tablefuncscan
|
|
* Determines and returns the cost of scanning a table function.
|
|
*
|
|
* 'baserel' is the relation to be scanned
|
|
* 'param_info' is the ParamPathInfo if this is a parameterized path, else NULL
|
|
*/
|
|
void
|
|
cost_tablefuncscan(Path *path, PlannerInfo *root,
|
|
RelOptInfo *baserel, ParamPathInfo *param_info)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
QualCost qpqual_cost;
|
|
Cost cpu_per_tuple;
|
|
RangeTblEntry *rte;
|
|
QualCost exprcost;
|
|
|
|
/* Should only be applied to base relations that are functions */
|
|
Assert(baserel->relid > 0);
|
|
rte = planner_rt_fetch(baserel->relid, root);
|
|
Assert(rte->rtekind == RTE_TABLEFUNC);
|
|
|
|
/* Mark the path with the correct row estimate */
|
|
if (param_info)
|
|
path->rows = param_info->ppi_rows;
|
|
else
|
|
path->rows = baserel->rows;
|
|
|
|
/*
|
|
* Estimate costs of executing the table func expression(s).
|
|
*
|
|
* XXX in principle we ought to charge tuplestore spill costs if the
|
|
* number of rows is large. However, given how phony our rowcount
|
|
* estimates for tablefuncs tend to be, there's not a lot of point in that
|
|
* refinement right now.
|
|
*/
|
|
cost_qual_eval_node(&exprcost, (Node *) rte->tablefunc, root);
|
|
|
|
startup_cost += exprcost.startup + exprcost.per_tuple;
|
|
|
|
/* Add scanning CPU costs */
|
|
get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
|
|
|
|
startup_cost += qpqual_cost.startup;
|
|
cpu_per_tuple = cpu_tuple_cost + qpqual_cost.per_tuple;
|
|
run_cost += cpu_per_tuple * baserel->tuples;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->pathtarget->cost.startup;
|
|
run_cost += path->pathtarget->cost.per_tuple * path->rows;
|
|
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_valuesscan
|
|
* Determines and returns the cost of scanning a VALUES RTE.
|
|
*
|
|
* 'baserel' is the relation to be scanned
|
|
* 'param_info' is the ParamPathInfo if this is a parameterized path, else NULL
|
|
*/
|
|
void
|
|
cost_valuesscan(Path *path, PlannerInfo *root,
|
|
RelOptInfo *baserel, ParamPathInfo *param_info)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
QualCost qpqual_cost;
|
|
Cost cpu_per_tuple;
|
|
|
|
/* Should only be applied to base relations that are values lists */
|
|
Assert(baserel->relid > 0);
|
|
Assert(baserel->rtekind == RTE_VALUES);
|
|
|
|
/* Mark the path with the correct row estimate */
|
|
if (param_info)
|
|
path->rows = param_info->ppi_rows;
|
|
else
|
|
path->rows = baserel->rows;
|
|
|
|
/*
|
|
* For now, estimate list evaluation cost at one operator eval per list
|
|
* (probably pretty bogus, but is it worth being smarter?)
|
|
*/
|
|
cpu_per_tuple = cpu_operator_cost;
|
|
|
|
/* Add scanning CPU costs */
|
|
get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
|
|
|
|
startup_cost += qpqual_cost.startup;
|
|
cpu_per_tuple += cpu_tuple_cost + qpqual_cost.per_tuple;
|
|
run_cost += cpu_per_tuple * baserel->tuples;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->pathtarget->cost.startup;
|
|
run_cost += path->pathtarget->cost.per_tuple * path->rows;
|
|
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_ctescan
|
|
* Determines and returns the cost of scanning a CTE RTE.
|
|
*
|
|
* Note: this is used for both self-reference and regular CTEs; the
|
|
* possible cost differences are below the threshold of what we could
|
|
* estimate accurately anyway. Note that the costs of evaluating the
|
|
* referenced CTE query are added into the final plan as initplan costs,
|
|
* and should NOT be counted here.
|
|
*/
|
|
void
|
|
cost_ctescan(Path *path, PlannerInfo *root,
|
|
RelOptInfo *baserel, ParamPathInfo *param_info)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
QualCost qpqual_cost;
|
|
Cost cpu_per_tuple;
|
|
|
|
/* Should only be applied to base relations that are CTEs */
|
|
Assert(baserel->relid > 0);
|
|
Assert(baserel->rtekind == RTE_CTE);
|
|
|
|
/* Mark the path with the correct row estimate */
|
|
if (param_info)
|
|
path->rows = param_info->ppi_rows;
|
|
else
|
|
path->rows = baserel->rows;
|
|
|
|
/* Charge one CPU tuple cost per row for tuplestore manipulation */
|
|
cpu_per_tuple = cpu_tuple_cost;
|
|
|
|
/* Add scanning CPU costs */
|
|
get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
|
|
|
|
startup_cost += qpqual_cost.startup;
|
|
cpu_per_tuple += cpu_tuple_cost + qpqual_cost.per_tuple;
|
|
run_cost += cpu_per_tuple * baserel->tuples;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->pathtarget->cost.startup;
|
|
run_cost += path->pathtarget->cost.per_tuple * path->rows;
|
|
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_namedtuplestorescan
|
|
* Determines and returns the cost of scanning a named tuplestore.
|
|
*/
|
|
void
|
|
cost_namedtuplestorescan(Path *path, PlannerInfo *root,
|
|
RelOptInfo *baserel, ParamPathInfo *param_info)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
QualCost qpqual_cost;
|
|
Cost cpu_per_tuple;
|
|
|
|
/* Should only be applied to base relations that are Tuplestores */
|
|
Assert(baserel->relid > 0);
|
|
Assert(baserel->rtekind == RTE_NAMEDTUPLESTORE);
|
|
|
|
/* Mark the path with the correct row estimate */
|
|
if (param_info)
|
|
path->rows = param_info->ppi_rows;
|
|
else
|
|
path->rows = baserel->rows;
|
|
|
|
/* Charge one CPU tuple cost per row for tuplestore manipulation */
|
|
cpu_per_tuple = cpu_tuple_cost;
|
|
|
|
/* Add scanning CPU costs */
|
|
get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
|
|
|
|
startup_cost += qpqual_cost.startup;
|
|
cpu_per_tuple += cpu_tuple_cost + qpqual_cost.per_tuple;
|
|
run_cost += cpu_per_tuple * baserel->tuples;
|
|
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_resultscan
|
|
* Determines and returns the cost of scanning an RTE_RESULT relation.
|
|
*/
|
|
void
|
|
cost_resultscan(Path *path, PlannerInfo *root,
|
|
RelOptInfo *baserel, ParamPathInfo *param_info)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
QualCost qpqual_cost;
|
|
Cost cpu_per_tuple;
|
|
|
|
/* Should only be applied to RTE_RESULT base relations */
|
|
Assert(baserel->relid > 0);
|
|
Assert(baserel->rtekind == RTE_RESULT);
|
|
|
|
/* Mark the path with the correct row estimate */
|
|
if (param_info)
|
|
path->rows = param_info->ppi_rows;
|
|
else
|
|
path->rows = baserel->rows;
|
|
|
|
/* We charge qual cost plus cpu_tuple_cost */
|
|
get_restriction_qual_cost(root, baserel, param_info, &qpqual_cost);
|
|
|
|
startup_cost += qpqual_cost.startup;
|
|
cpu_per_tuple = cpu_tuple_cost + qpqual_cost.per_tuple;
|
|
run_cost += cpu_per_tuple * baserel->tuples;
|
|
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_recursive_union
|
|
* Determines and returns the cost of performing a recursive union,
|
|
* and also the estimated output size.
|
|
*
|
|
* We are given Paths for the nonrecursive and recursive terms.
|
|
*/
|
|
void
|
|
cost_recursive_union(Path *runion, Path *nrterm, Path *rterm)
|
|
{
|
|
Cost startup_cost;
|
|
Cost total_cost;
|
|
double total_rows;
|
|
|
|
/* We probably have decent estimates for the non-recursive term */
|
|
startup_cost = nrterm->startup_cost;
|
|
total_cost = nrterm->total_cost;
|
|
total_rows = nrterm->rows;
|
|
|
|
/*
|
|
* We arbitrarily assume that about 10 recursive iterations will be
|
|
* needed, and that we've managed to get a good fix on the cost and output
|
|
* size of each one of them. These are mighty shaky assumptions but it's
|
|
* hard to see how to do better.
|
|
*/
|
|
total_cost += 10 * rterm->total_cost;
|
|
total_rows += 10 * rterm->rows;
|
|
|
|
/*
|
|
* Also charge cpu_tuple_cost per row to account for the costs of
|
|
* manipulating the tuplestores. (We don't worry about possible
|
|
* spill-to-disk costs.)
|
|
*/
|
|
total_cost += cpu_tuple_cost * total_rows;
|
|
|
|
runion->startup_cost = startup_cost;
|
|
runion->total_cost = total_cost;
|
|
runion->rows = total_rows;
|
|
runion->pathtarget->width = Max(nrterm->pathtarget->width,
|
|
rterm->pathtarget->width);
|
|
}
|
|
|
|
/*
|
|
* cost_tuplesort
|
|
* Determines and returns the cost of sorting a relation using tuplesort,
|
|
* not including the cost of reading the input data.
|
|
*
|
|
* If the total volume of data to sort is less than sort_mem, we will do
|
|
* an in-memory sort, which requires no I/O and about t*log2(t) tuple
|
|
* comparisons for t tuples.
|
|
*
|
|
* If the total volume exceeds sort_mem, we switch to a tape-style merge
|
|
* algorithm. There will still be about t*log2(t) tuple comparisons in
|
|
* total, but we will also need to write and read each tuple once per
|
|
* merge pass. We expect about ceil(logM(r)) merge passes where r is the
|
|
* number of initial runs formed and M is the merge order used by tuplesort.c.
|
|
* Since the average initial run should be about sort_mem, we have
|
|
* disk traffic = 2 * relsize * ceil(logM(p / sort_mem))
|
|
* cpu = comparison_cost * t * log2(t)
|
|
*
|
|
* If the sort is bounded (i.e., only the first k result tuples are needed)
|
|
* and k tuples can fit into sort_mem, we use a heap method that keeps only
|
|
* k tuples in the heap; this will require about t*log2(k) tuple comparisons.
|
|
*
|
|
* The disk traffic is assumed to be 3/4ths sequential and 1/4th random
|
|
* accesses (XXX can't we refine that guess?)
|
|
*
|
|
* By default, we charge two operator evals per tuple comparison, which should
|
|
* be in the right ballpark in most cases. The caller can tweak this by
|
|
* specifying nonzero comparison_cost; typically that's used for any extra
|
|
* work that has to be done to prepare the inputs to the comparison operators.
|
|
*
|
|
* 'tuples' is the number of tuples in the relation
|
|
* 'width' is the average tuple width in bytes
|
|
* 'comparison_cost' is the extra cost per comparison, if any
|
|
* 'sort_mem' is the number of kilobytes of work memory allowed for the sort
|
|
* 'limit_tuples' is the bound on the number of output tuples; -1 if no bound
|
|
*/
|
|
static void
|
|
cost_tuplesort(Cost *startup_cost, Cost *run_cost,
|
|
double tuples, int width,
|
|
Cost comparison_cost, int sort_mem,
|
|
double limit_tuples)
|
|
{
|
|
double input_bytes = relation_byte_size(tuples, width);
|
|
double output_bytes;
|
|
double output_tuples;
|
|
long sort_mem_bytes = sort_mem * 1024L;
|
|
|
|
/*
|
|
* We want to be sure the cost of a sort is never estimated as zero, even
|
|
* if passed-in tuple count is zero. Besides, mustn't do log(0)...
|
|
*/
|
|
if (tuples < 2.0)
|
|
tuples = 2.0;
|
|
|
|
/* Include the default cost-per-comparison */
|
|
comparison_cost += 2.0 * cpu_operator_cost;
|
|
|
|
/* Do we have a useful LIMIT? */
|
|
if (limit_tuples > 0 && limit_tuples < tuples)
|
|
{
|
|
output_tuples = limit_tuples;
|
|
output_bytes = relation_byte_size(output_tuples, width);
|
|
}
|
|
else
|
|
{
|
|
output_tuples = tuples;
|
|
output_bytes = input_bytes;
|
|
}
|
|
|
|
if (output_bytes > sort_mem_bytes)
|
|
{
|
|
/*
|
|
* We'll have to use a disk-based sort of all the tuples
|
|
*/
|
|
double npages = ceil(input_bytes / BLCKSZ);
|
|
double nruns = input_bytes / sort_mem_bytes;
|
|
double mergeorder = tuplesort_merge_order(sort_mem_bytes);
|
|
double log_runs;
|
|
double npageaccesses;
|
|
|
|
/*
|
|
* CPU costs
|
|
*
|
|
* Assume about N log2 N comparisons
|
|
*/
|
|
*startup_cost = comparison_cost * tuples * LOG2(tuples);
|
|
|
|
/* Disk costs */
|
|
|
|
/* Compute logM(r) as log(r) / log(M) */
|
|
if (nruns > mergeorder)
|
|
log_runs = ceil(log(nruns) / log(mergeorder));
|
|
else
|
|
log_runs = 1.0;
|
|
npageaccesses = 2.0 * npages * log_runs;
|
|
/* Assume 3/4ths of accesses are sequential, 1/4th are not */
|
|
*startup_cost += npageaccesses *
|
|
(seq_page_cost * 0.75 + random_page_cost * 0.25);
|
|
}
|
|
else if (tuples > 2 * output_tuples || input_bytes > sort_mem_bytes)
|
|
{
|
|
/*
|
|
* We'll use a bounded heap-sort keeping just K tuples in memory, for
|
|
* a total number of tuple comparisons of N log2 K; but the constant
|
|
* factor is a bit higher than for quicksort. Tweak it so that the
|
|
* cost curve is continuous at the crossover point.
|
|
*/
|
|
*startup_cost = comparison_cost * tuples * LOG2(2.0 * output_tuples);
|
|
}
|
|
else
|
|
{
|
|
/* We'll use plain quicksort on all the input tuples */
|
|
*startup_cost = comparison_cost * tuples * LOG2(tuples);
|
|
}
|
|
|
|
/*
|
|
* Also charge a small amount (arbitrarily set equal to operator cost) per
|
|
* extracted tuple. We don't charge cpu_tuple_cost because a Sort node
|
|
* doesn't do qual-checking or projection, so it has less overhead than
|
|
* most plan nodes. Note it's correct to use tuples not output_tuples
|
|
* here --- the upper LIMIT will pro-rate the run cost so we'd be double
|
|
* counting the LIMIT otherwise.
|
|
*/
|
|
*run_cost = cpu_operator_cost * tuples;
|
|
}
|
|
|
|
/*
|
|
* cost_incremental_sort
|
|
* Determines and returns the cost of sorting a relation incrementally, when
|
|
* the input path is presorted by a prefix of the pathkeys.
|
|
*
|
|
* 'presorted_keys' is the number of leading pathkeys by which the input path
|
|
* is sorted.
|
|
*
|
|
* We estimate the number of groups into which the relation is divided by the
|
|
* leading pathkeys, and then calculate the cost of sorting a single group
|
|
* with tuplesort using cost_tuplesort().
|
|
*/
|
|
void
|
|
cost_incremental_sort(Path *path,
|
|
PlannerInfo *root, List *pathkeys, int presorted_keys,
|
|
Cost input_startup_cost, Cost input_total_cost,
|
|
double input_tuples, int width, Cost comparison_cost, int sort_mem,
|
|
double limit_tuples)
|
|
{
|
|
Cost startup_cost,
|
|
run_cost,
|
|
input_run_cost = input_total_cost - input_startup_cost;
|
|
double group_tuples,
|
|
input_groups;
|
|
Cost group_startup_cost,
|
|
group_run_cost,
|
|
group_input_run_cost;
|
|
List *presortedExprs = NIL;
|
|
ListCell *l;
|
|
bool unknown_varno = false;
|
|
|
|
Assert(presorted_keys > 0 && presorted_keys < list_length(pathkeys));
|
|
|
|
/*
|
|
* We want to be sure the cost of a sort is never estimated as zero, even
|
|
* if passed-in tuple count is zero. Besides, mustn't do log(0)...
|
|
*/
|
|
if (input_tuples < 2.0)
|
|
input_tuples = 2.0;
|
|
|
|
/* Default estimate of number of groups, capped to one group per row. */
|
|
input_groups = Min(input_tuples, DEFAULT_NUM_DISTINCT);
|
|
|
|
/*
|
|
* Extract presorted keys as list of expressions.
|
|
*
|
|
* We need to be careful about Vars containing "varno 0" which might have
|
|
* been introduced by generate_append_tlist, which would confuse
|
|
* estimate_num_groups (in fact it'd fail for such expressions). See
|
|
* recurse_set_operations which has to deal with the same issue.
|
|
*
|
|
* Unlike recurse_set_operations we can't access the original target list
|
|
* here, and even if we could it's not very clear how useful would that be
|
|
* for a set operation combining multiple tables. So we simply detect if
|
|
* there are any expressions with "varno 0" and use the default
|
|
* DEFAULT_NUM_DISTINCT in that case.
|
|
*
|
|
* We might also use either 1.0 (a single group) or input_tuples (each row
|
|
* being a separate group), pretty much the worst and best case for
|
|
* incremental sort. But those are extreme cases and using something in
|
|
* between seems reasonable. Furthermore, generate_append_tlist is used
|
|
* for set operations, which are likely to produce mostly unique output
|
|
* anyway - from that standpoint the DEFAULT_NUM_DISTINCT is defensive
|
|
* while maintaining lower startup cost.
|
|
*/
|
|
foreach(l, pathkeys)
|
|
{
|
|
PathKey *key = (PathKey *) lfirst(l);
|
|
EquivalenceMember *member = (EquivalenceMember *)
|
|
linitial(key->pk_eclass->ec_members);
|
|
|
|
/*
|
|
* Check if the expression contains Var with "varno 0" so that we
|
|
* don't call estimate_num_groups in that case.
|
|
*/
|
|
if (bms_is_member(0, pull_varnos(root, (Node *) member->em_expr)))
|
|
{
|
|
unknown_varno = true;
|
|
break;
|
|
}
|
|
|
|
/* expression not containing any Vars with "varno 0" */
|
|
presortedExprs = lappend(presortedExprs, member->em_expr);
|
|
|
|
if (foreach_current_index(l) + 1 >= presorted_keys)
|
|
break;
|
|
}
|
|
|
|
/* Estimate the number of groups with equal presorted keys. */
|
|
if (!unknown_varno)
|
|
input_groups = estimate_num_groups(root, presortedExprs, input_tuples,
|
|
NULL, NULL);
|
|
|
|
group_tuples = input_tuples / input_groups;
|
|
group_input_run_cost = input_run_cost / input_groups;
|
|
|
|
/*
|
|
* Estimate the average cost of sorting of one group where presorted keys
|
|
* are equal.
|
|
*/
|
|
cost_tuplesort(&group_startup_cost, &group_run_cost,
|
|
group_tuples, width, comparison_cost, sort_mem,
|
|
limit_tuples);
|
|
|
|
/*
|
|
* Startup cost of incremental sort is the startup cost of its first group
|
|
* plus the cost of its input.
|
|
*/
|
|
startup_cost = group_startup_cost + input_startup_cost +
|
|
group_input_run_cost;
|
|
|
|
/*
|
|
* After we started producing tuples from the first group, the cost of
|
|
* producing all the tuples is given by the cost to finish processing this
|
|
* group, plus the total cost to process the remaining groups, plus the
|
|
* remaining cost of input.
|
|
*/
|
|
run_cost = group_run_cost + (group_run_cost + group_startup_cost) *
|
|
(input_groups - 1) + group_input_run_cost * (input_groups - 1);
|
|
|
|
/*
|
|
* Incremental sort adds some overhead by itself. Firstly, it has to
|
|
* detect the sort groups. This is roughly equal to one extra copy and
|
|
* comparison per tuple.
|
|
*/
|
|
run_cost += (cpu_tuple_cost + comparison_cost) * input_tuples;
|
|
|
|
/*
|
|
* Additionally, we charge double cpu_tuple_cost for each input group to
|
|
* account for the tuplesort_reset that's performed after each group.
|
|
*/
|
|
run_cost += 2.0 * cpu_tuple_cost * input_groups;
|
|
|
|
path->rows = input_tuples;
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_sort
|
|
* Determines and returns the cost of sorting a relation, including
|
|
* the cost of reading the input data.
|
|
*
|
|
* NOTE: some callers currently pass NIL for pathkeys because they
|
|
* can't conveniently supply the sort keys. Since this routine doesn't
|
|
* currently do anything with pathkeys anyway, that doesn't matter...
|
|
* but if it ever does, it should react gracefully to lack of key data.
|
|
* (Actually, the thing we'd most likely be interested in is just the number
|
|
* of sort keys, which all callers *could* supply.)
|
|
*/
|
|
void
|
|
cost_sort(Path *path, PlannerInfo *root,
|
|
List *pathkeys, Cost input_cost, double tuples, int width,
|
|
Cost comparison_cost, int sort_mem,
|
|
double limit_tuples)
|
|
|
|
{
|
|
Cost startup_cost;
|
|
Cost run_cost;
|
|
|
|
cost_tuplesort(&startup_cost, &run_cost,
|
|
tuples, width,
|
|
comparison_cost, sort_mem,
|
|
limit_tuples);
|
|
|
|
if (!enable_sort)
|
|
startup_cost += disable_cost;
|
|
|
|
startup_cost += input_cost;
|
|
|
|
path->rows = tuples;
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* append_nonpartial_cost
|
|
* Estimate the cost of the non-partial paths in a Parallel Append.
|
|
* The non-partial paths are assumed to be the first "numpaths" paths
|
|
* from the subpaths list, and to be in order of decreasing cost.
|
|
*/
|
|
static Cost
|
|
append_nonpartial_cost(List *subpaths, int numpaths, int parallel_workers)
|
|
{
|
|
Cost *costarr;
|
|
int arrlen;
|
|
ListCell *l;
|
|
ListCell *cell;
|
|
int path_index;
|
|
int min_index;
|
|
int max_index;
|
|
|
|
if (numpaths == 0)
|
|
return 0;
|
|
|
|
/*
|
|
* Array length is number of workers or number of relevant paths,
|
|
* whichever is less.
|
|
*/
|
|
arrlen = Min(parallel_workers, numpaths);
|
|
costarr = (Cost *) palloc(sizeof(Cost) * arrlen);
|
|
|
|
/* The first few paths will each be claimed by a different worker. */
|
|
path_index = 0;
|
|
foreach(cell, subpaths)
|
|
{
|
|
Path *subpath = (Path *) lfirst(cell);
|
|
|
|
if (path_index == arrlen)
|
|
break;
|
|
costarr[path_index++] = subpath->total_cost;
|
|
}
|
|
|
|
/*
|
|
* Since subpaths are sorted by decreasing cost, the last one will have
|
|
* the minimum cost.
|
|
*/
|
|
min_index = arrlen - 1;
|
|
|
|
/*
|
|
* For each of the remaining subpaths, add its cost to the array element
|
|
* with minimum cost.
|
|
*/
|
|
for_each_cell(l, subpaths, cell)
|
|
{
|
|
Path *subpath = (Path *) lfirst(l);
|
|
|
|
/* Consider only the non-partial paths */
|
|
if (path_index++ == numpaths)
|
|
break;
|
|
|
|
costarr[min_index] += subpath->total_cost;
|
|
|
|
/* Update the new min cost array index */
|
|
min_index = 0;
|
|
for (int i = 0; i < arrlen; i++)
|
|
{
|
|
if (costarr[i] < costarr[min_index])
|
|
min_index = i;
|
|
}
|
|
}
|
|
|
|
/* Return the highest cost from the array */
|
|
max_index = 0;
|
|
for (int i = 0; i < arrlen; i++)
|
|
{
|
|
if (costarr[i] > costarr[max_index])
|
|
max_index = i;
|
|
}
|
|
|
|
return costarr[max_index];
|
|
}
|
|
|
|
/*
|
|
* cost_append
|
|
* Determines and returns the cost of an Append node.
|
|
*/
|
|
void
|
|
cost_append(AppendPath *apath)
|
|
{
|
|
ListCell *l;
|
|
|
|
apath->path.startup_cost = 0;
|
|
apath->path.total_cost = 0;
|
|
apath->path.rows = 0;
|
|
|
|
if (apath->subpaths == NIL)
|
|
return;
|
|
|
|
if (!apath->path.parallel_aware)
|
|
{
|
|
List *pathkeys = apath->path.pathkeys;
|
|
|
|
if (pathkeys == NIL)
|
|
{
|
|
Path *firstsubpath = (Path *) linitial(apath->subpaths);
|
|
|
|
/*
|
|
* For an unordered, non-parallel-aware Append we take the startup
|
|
* cost as the startup cost of the first subpath.
|
|
*/
|
|
apath->path.startup_cost = firstsubpath->startup_cost;
|
|
|
|
/* Compute rows and costs as sums of subplan rows and costs. */
|
|
foreach(l, apath->subpaths)
|
|
{
|
|
Path *subpath = (Path *) lfirst(l);
|
|
|
|
apath->path.rows += subpath->rows;
|
|
apath->path.total_cost += subpath->total_cost;
|
|
}
|
|
}
|
|
else
|
|
{
|
|
/*
|
|
* For an ordered, non-parallel-aware Append we take the startup
|
|
* cost as the sum of the subpath startup costs. This ensures
|
|
* that we don't underestimate the startup cost when a query's
|
|
* LIMIT is such that several of the children have to be run to
|
|
* satisfy it. This might be overkill --- another plausible hack
|
|
* would be to take the Append's startup cost as the maximum of
|
|
* the child startup costs. But we don't want to risk believing
|
|
* that an ORDER BY LIMIT query can be satisfied at small cost
|
|
* when the first child has small startup cost but later ones
|
|
* don't. (If we had the ability to deal with nonlinear cost
|
|
* interpolation for partial retrievals, we would not need to be
|
|
* so conservative about this.)
|
|
*
|
|
* This case is also different from the above in that we have to
|
|
* account for possibly injecting sorts into subpaths that aren't
|
|
* natively ordered.
|
|
*/
|
|
foreach(l, apath->subpaths)
|
|
{
|
|
Path *subpath = (Path *) lfirst(l);
|
|
Path sort_path; /* dummy for result of cost_sort */
|
|
|
|
if (!pathkeys_contained_in(pathkeys, subpath->pathkeys))
|
|
{
|
|
/*
|
|
* We'll need to insert a Sort node, so include costs for
|
|
* that. We can use the parent's LIMIT if any, since we
|
|
* certainly won't pull more than that many tuples from
|
|
* any child.
|
|
*/
|
|
cost_sort(&sort_path,
|
|
NULL, /* doesn't currently need root */
|
|
pathkeys,
|
|
subpath->total_cost,
|
|
subpath->rows,
|
|
subpath->pathtarget->width,
|
|
0.0,
|
|
work_mem,
|
|
apath->limit_tuples);
|
|
subpath = &sort_path;
|
|
}
|
|
|
|
apath->path.rows += subpath->rows;
|
|
apath->path.startup_cost += subpath->startup_cost;
|
|
apath->path.total_cost += subpath->total_cost;
|
|
}
|
|
}
|
|
}
|
|
else /* parallel-aware */
|
|
{
|
|
int i = 0;
|
|
double parallel_divisor = get_parallel_divisor(&apath->path);
|
|
|
|
/* Parallel-aware Append never produces ordered output. */
|
|
Assert(apath->path.pathkeys == NIL);
|
|
|
|
/* Calculate startup cost. */
|
|
foreach(l, apath->subpaths)
|
|
{
|
|
Path *subpath = (Path *) lfirst(l);
|
|
|
|
/*
|
|
* Append will start returning tuples when the child node having
|
|
* lowest startup cost is done setting up. We consider only the
|
|
* first few subplans that immediately get a worker assigned.
|
|
*/
|
|
if (i == 0)
|
|
apath->path.startup_cost = subpath->startup_cost;
|
|
else if (i < apath->path.parallel_workers)
|
|
apath->path.startup_cost = Min(apath->path.startup_cost,
|
|
subpath->startup_cost);
|
|
|
|
/*
|
|
* Apply parallel divisor to subpaths. Scale the number of rows
|
|
* for each partial subpath based on the ratio of the parallel
|
|
* divisor originally used for the subpath to the one we adopted.
|
|
* Also add the cost of partial paths to the total cost, but
|
|
* ignore non-partial paths for now.
|
|
*/
|
|
if (i < apath->first_partial_path)
|
|
apath->path.rows += subpath->rows / parallel_divisor;
|
|
else
|
|
{
|
|
double subpath_parallel_divisor;
|
|
|
|
subpath_parallel_divisor = get_parallel_divisor(subpath);
|
|
apath->path.rows += subpath->rows * (subpath_parallel_divisor /
|
|
parallel_divisor);
|
|
apath->path.total_cost += subpath->total_cost;
|
|
}
|
|
|
|
apath->path.rows = clamp_row_est(apath->path.rows);
|
|
|
|
i++;
|
|
}
|
|
|
|
/* Add cost for non-partial subpaths. */
|
|
apath->path.total_cost +=
|
|
append_nonpartial_cost(apath->subpaths,
|
|
apath->first_partial_path,
|
|
apath->path.parallel_workers);
|
|
}
|
|
|
|
/*
|
|
* Although Append does not do any selection or projection, it's not free;
|
|
* add a small per-tuple overhead.
|
|
*/
|
|
apath->path.total_cost +=
|
|
cpu_tuple_cost * APPEND_CPU_COST_MULTIPLIER * apath->path.rows;
|
|
}
|
|
|
|
/*
|
|
* cost_merge_append
|
|
* Determines and returns the cost of a MergeAppend node.
|
|
*
|
|
* MergeAppend merges several pre-sorted input streams, using a heap that
|
|
* at any given instant holds the next tuple from each stream. If there
|
|
* are N streams, we need about N*log2(N) tuple comparisons to construct
|
|
* the heap at startup, and then for each output tuple, about log2(N)
|
|
* comparisons to replace the top entry.
|
|
*
|
|
* (The effective value of N will drop once some of the input streams are
|
|
* exhausted, but it seems unlikely to be worth trying to account for that.)
|
|
*
|
|
* The heap is never spilled to disk, since we assume N is not very large.
|
|
* So this is much simpler than cost_sort.
|
|
*
|
|
* As in cost_sort, we charge two operator evals per tuple comparison.
|
|
*
|
|
* 'pathkeys' is a list of sort keys
|
|
* 'n_streams' is the number of input streams
|
|
* 'input_startup_cost' is the sum of the input streams' startup costs
|
|
* 'input_total_cost' is the sum of the input streams' total costs
|
|
* 'tuples' is the number of tuples in all the streams
|
|
*/
|
|
void
|
|
cost_merge_append(Path *path, PlannerInfo *root,
|
|
List *pathkeys, int n_streams,
|
|
Cost input_startup_cost, Cost input_total_cost,
|
|
double tuples)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
Cost comparison_cost;
|
|
double N;
|
|
double logN;
|
|
|
|
/*
|
|
* Avoid log(0)...
|
|
*/
|
|
N = (n_streams < 2) ? 2.0 : (double) n_streams;
|
|
logN = LOG2(N);
|
|
|
|
/* Assumed cost per tuple comparison */
|
|
comparison_cost = 2.0 * cpu_operator_cost;
|
|
|
|
/* Heap creation cost */
|
|
startup_cost += comparison_cost * N * logN;
|
|
|
|
/* Per-tuple heap maintenance cost */
|
|
run_cost += tuples * comparison_cost * logN;
|
|
|
|
/*
|
|
* Although MergeAppend does not do any selection or projection, it's not
|
|
* free; add a small per-tuple overhead.
|
|
*/
|
|
run_cost += cpu_tuple_cost * APPEND_CPU_COST_MULTIPLIER * tuples;
|
|
|
|
path->startup_cost = startup_cost + input_startup_cost;
|
|
path->total_cost = startup_cost + run_cost + input_total_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_material
|
|
* Determines and returns the cost of materializing a relation, including
|
|
* the cost of reading the input data.
|
|
*
|
|
* If the total volume of data to materialize exceeds work_mem, we will need
|
|
* to write it to disk, so the cost is much higher in that case.
|
|
*
|
|
* Note that here we are estimating the costs for the first scan of the
|
|
* relation, so the materialization is all overhead --- any savings will
|
|
* occur only on rescan, which is estimated in cost_rescan.
|
|
*/
|
|
void
|
|
cost_material(Path *path,
|
|
Cost input_startup_cost, Cost input_total_cost,
|
|
double tuples, int width)
|
|
{
|
|
Cost startup_cost = input_startup_cost;
|
|
Cost run_cost = input_total_cost - input_startup_cost;
|
|
double nbytes = relation_byte_size(tuples, width);
|
|
long work_mem_bytes = work_mem * 1024L;
|
|
|
|
path->rows = tuples;
|
|
|
|
/*
|
|
* Whether spilling or not, charge 2x cpu_operator_cost per tuple to
|
|
* reflect bookkeeping overhead. (This rate must be more than what
|
|
* cost_rescan charges for materialize, ie, cpu_operator_cost per tuple;
|
|
* if it is exactly the same then there will be a cost tie between
|
|
* nestloop with A outer, materialized B inner and nestloop with B outer,
|
|
* materialized A inner. The extra cost ensures we'll prefer
|
|
* materializing the smaller rel.) Note that this is normally a good deal
|
|
* less than cpu_tuple_cost; which is OK because a Material plan node
|
|
* doesn't do qual-checking or projection, so it's got less overhead than
|
|
* most plan nodes.
|
|
*/
|
|
run_cost += 2 * cpu_operator_cost * tuples;
|
|
|
|
/*
|
|
* If we will spill to disk, charge at the rate of seq_page_cost per page.
|
|
* This cost is assumed to be evenly spread through the plan run phase,
|
|
* which isn't exactly accurate but our cost model doesn't allow for
|
|
* nonuniform costs within the run phase.
|
|
*/
|
|
if (nbytes > work_mem_bytes)
|
|
{
|
|
double npages = ceil(nbytes / BLCKSZ);
|
|
|
|
run_cost += seq_page_cost * npages;
|
|
}
|
|
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_memoize_rescan
|
|
* Determines the estimated cost of rescanning a Memoize node.
|
|
*
|
|
* In order to estimate this, we must gain knowledge of how often we expect to
|
|
* be called and how many distinct sets of parameters we are likely to be
|
|
* called with. If we expect a good cache hit ratio, then we can set our
|
|
* costs to account for that hit ratio, plus a little bit of cost for the
|
|
* caching itself. Caching will not work out well if we expect to be called
|
|
* with too many distinct parameter values. The worst-case here is that we
|
|
* never see any parameter value twice, in which case we'd never get a cache
|
|
* hit and caching would be a complete waste of effort.
|
|
*/
|
|
static void
|
|
cost_memoize_rescan(PlannerInfo *root, MemoizePath *mpath,
|
|
Cost *rescan_startup_cost, Cost *rescan_total_cost)
|
|
{
|
|
EstimationInfo estinfo;
|
|
ListCell *lc;
|
|
Cost input_startup_cost = mpath->subpath->startup_cost;
|
|
Cost input_total_cost = mpath->subpath->total_cost;
|
|
double tuples = mpath->subpath->rows;
|
|
double calls = mpath->calls;
|
|
int width = mpath->subpath->pathtarget->width;
|
|
|
|
double hash_mem_bytes;
|
|
double est_entry_bytes;
|
|
double est_cache_entries;
|
|
double ndistinct;
|
|
double evict_ratio;
|
|
double hit_ratio;
|
|
Cost startup_cost;
|
|
Cost total_cost;
|
|
|
|
/* available cache space */
|
|
hash_mem_bytes = get_hash_memory_limit();
|
|
|
|
/*
|
|
* Set the number of bytes each cache entry should consume in the cache.
|
|
* To provide us with better estimations on how many cache entries we can
|
|
* store at once, we make a call to the executor here to ask it what
|
|
* memory overheads there are for a single cache entry.
|
|
*/
|
|
est_entry_bytes = relation_byte_size(tuples, width) +
|
|
ExecEstimateCacheEntryOverheadBytes(tuples);
|
|
|
|
/* include the estimated width for the cache keys */
|
|
foreach(lc, mpath->param_exprs)
|
|
est_entry_bytes += get_expr_width(root, (Node *) lfirst(lc));
|
|
|
|
/* estimate on the upper limit of cache entries we can hold at once */
|
|
est_cache_entries = floor(hash_mem_bytes / est_entry_bytes);
|
|
|
|
/* estimate on the distinct number of parameter values */
|
|
ndistinct = estimate_num_groups(root, mpath->param_exprs, calls, NULL,
|
|
&estinfo);
|
|
|
|
/*
|
|
* When the estimation fell back on using a default value, it's a bit too
|
|
* risky to assume that it's ok to use a Memoize node. The use of a
|
|
* default could cause us to use a Memoize node when it's really
|
|
* inappropriate to do so. If we see that this has been done, then we'll
|
|
* assume that every call will have unique parameters, which will almost
|
|
* certainly mean a MemoizePath will never survive add_path().
|
|
*/
|
|
if ((estinfo.flags & SELFLAG_USED_DEFAULT) != 0)
|
|
ndistinct = calls;
|
|
|
|
/*
|
|
* Since we've already estimated the maximum number of entries we can
|
|
* store at once and know the estimated number of distinct values we'll be
|
|
* called with, we'll take this opportunity to set the path's est_entries.
|
|
* This will ultimately determine the hash table size that the executor
|
|
* will use. If we leave this at zero, the executor will just choose the
|
|
* size itself. Really this is not the right place to do this, but it's
|
|
* convenient since everything is already calculated.
|
|
*/
|
|
mpath->est_entries = Min(Min(ndistinct, est_cache_entries),
|
|
PG_UINT32_MAX);
|
|
|
|
/*
|
|
* When the number of distinct parameter values is above the amount we can
|
|
* store in the cache, then we'll have to evict some entries from the
|
|
* cache. This is not free. Here we estimate how often we'll incur the
|
|
* cost of that eviction.
|
|
*/
|
|
evict_ratio = 1.0 - Min(est_cache_entries, ndistinct) / ndistinct;
|
|
|
|
/*
|
|
* In order to estimate how costly a single scan will be, we need to
|
|
* attempt to estimate what the cache hit ratio will be. To do that we
|
|
* must look at how many scans are estimated in total for this node and
|
|
* how many of those scans we expect to get a cache hit.
|
|
*/
|
|
hit_ratio = ((calls - ndistinct) / calls) *
|
|
(est_cache_entries / Max(ndistinct, est_cache_entries));
|
|
|
|
Assert(hit_ratio >= 0 && hit_ratio <= 1.0);
|
|
|
|
/*
|
|
* Set the total_cost accounting for the expected cache hit ratio. We
|
|
* also add on a cpu_operator_cost to account for a cache lookup. This
|
|
* will happen regardless of whether it's a cache hit or not.
|
|
*/
|
|
total_cost = input_total_cost * (1.0 - hit_ratio) + cpu_operator_cost;
|
|
|
|
/* Now adjust the total cost to account for cache evictions */
|
|
|
|
/* Charge a cpu_tuple_cost for evicting the actual cache entry */
|
|
total_cost += cpu_tuple_cost * evict_ratio;
|
|
|
|
/*
|
|
* Charge a 10th of cpu_operator_cost to evict every tuple in that entry.
|
|
* The per-tuple eviction is really just a pfree, so charging a whole
|
|
* cpu_operator_cost seems a little excessive.
|
|
*/
|
|
total_cost += cpu_operator_cost / 10.0 * evict_ratio * tuples;
|
|
|
|
/*
|
|
* Now adjust for storing things in the cache, since that's not free
|
|
* either. Everything must go in the cache. We don't proportion this
|
|
* over any ratio, just apply it once for the scan. We charge a
|
|
* cpu_tuple_cost for the creation of the cache entry and also a
|
|
* cpu_operator_cost for each tuple we expect to cache.
|
|
*/
|
|
total_cost += cpu_tuple_cost + cpu_operator_cost * tuples;
|
|
|
|
/*
|
|
* Getting the first row must be also be proportioned according to the
|
|
* expected cache hit ratio.
|
|
*/
|
|
startup_cost = input_startup_cost * (1.0 - hit_ratio);
|
|
|
|
/*
|
|
* Additionally we charge a cpu_tuple_cost to account for cache lookups,
|
|
* which we'll do regardless of whether it was a cache hit or not.
|
|
*/
|
|
startup_cost += cpu_tuple_cost;
|
|
|
|
*rescan_startup_cost = startup_cost;
|
|
*rescan_total_cost = total_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_agg
|
|
* Determines and returns the cost of performing an Agg plan node,
|
|
* including the cost of its input.
|
|
*
|
|
* aggcosts can be NULL when there are no actual aggregate functions (i.e.,
|
|
* we are using a hashed Agg node just to do grouping).
|
|
*
|
|
* Note: when aggstrategy == AGG_SORTED, caller must ensure that input costs
|
|
* are for appropriately-sorted input.
|
|
*/
|
|
void
|
|
cost_agg(Path *path, PlannerInfo *root,
|
|
AggStrategy aggstrategy, const AggClauseCosts *aggcosts,
|
|
int numGroupCols, double numGroups,
|
|
List *quals,
|
|
Cost input_startup_cost, Cost input_total_cost,
|
|
double input_tuples, double input_width)
|
|
{
|
|
double output_tuples;
|
|
Cost startup_cost;
|
|
Cost total_cost;
|
|
AggClauseCosts dummy_aggcosts;
|
|
|
|
/* Use all-zero per-aggregate costs if NULL is passed */
|
|
if (aggcosts == NULL)
|
|
{
|
|
Assert(aggstrategy == AGG_HASHED);
|
|
MemSet(&dummy_aggcosts, 0, sizeof(AggClauseCosts));
|
|
aggcosts = &dummy_aggcosts;
|
|
}
|
|
|
|
/*
|
|
* The transCost.per_tuple component of aggcosts should be charged once
|
|
* per input tuple, corresponding to the costs of evaluating the aggregate
|
|
* transfns and their input expressions. The finalCost.per_tuple component
|
|
* is charged once per output tuple, corresponding to the costs of
|
|
* evaluating the finalfns. Startup costs are of course charged but once.
|
|
*
|
|
* If we are grouping, we charge an additional cpu_operator_cost per
|
|
* grouping column per input tuple for grouping comparisons.
|
|
*
|
|
* We will produce a single output tuple if not grouping, and a tuple per
|
|
* group otherwise. We charge cpu_tuple_cost for each output tuple.
|
|
*
|
|
* Note: in this cost model, AGG_SORTED and AGG_HASHED have exactly the
|
|
* same total CPU cost, but AGG_SORTED has lower startup cost. If the
|
|
* input path is already sorted appropriately, AGG_SORTED should be
|
|
* preferred (since it has no risk of memory overflow). This will happen
|
|
* as long as the computed total costs are indeed exactly equal --- but if
|
|
* there's roundoff error we might do the wrong thing. So be sure that
|
|
* the computations below form the same intermediate values in the same
|
|
* order.
|
|
*/
|
|
if (aggstrategy == AGG_PLAIN)
|
|
{
|
|
startup_cost = input_total_cost;
|
|
startup_cost += aggcosts->transCost.startup;
|
|
startup_cost += aggcosts->transCost.per_tuple * input_tuples;
|
|
startup_cost += aggcosts->finalCost.startup;
|
|
startup_cost += aggcosts->finalCost.per_tuple;
|
|
/* we aren't grouping */
|
|
total_cost = startup_cost + cpu_tuple_cost;
|
|
output_tuples = 1;
|
|
}
|
|
else if (aggstrategy == AGG_SORTED || aggstrategy == AGG_MIXED)
|
|
{
|
|
/* Here we are able to deliver output on-the-fly */
|
|
startup_cost = input_startup_cost;
|
|
total_cost = input_total_cost;
|
|
if (aggstrategy == AGG_MIXED && !enable_hashagg)
|
|
{
|
|
startup_cost += disable_cost;
|
|
total_cost += disable_cost;
|
|
}
|
|
/* calcs phrased this way to match HASHED case, see note above */
|
|
total_cost += aggcosts->transCost.startup;
|
|
total_cost += aggcosts->transCost.per_tuple * input_tuples;
|
|
total_cost += (cpu_operator_cost * numGroupCols) * input_tuples;
|
|
total_cost += aggcosts->finalCost.startup;
|
|
total_cost += aggcosts->finalCost.per_tuple * numGroups;
|
|
total_cost += cpu_tuple_cost * numGroups;
|
|
output_tuples = numGroups;
|
|
}
|
|
else
|
|
{
|
|
/* must be AGG_HASHED */
|
|
startup_cost = input_total_cost;
|
|
if (!enable_hashagg)
|
|
startup_cost += disable_cost;
|
|
startup_cost += aggcosts->transCost.startup;
|
|
startup_cost += aggcosts->transCost.per_tuple * input_tuples;
|
|
/* cost of computing hash value */
|
|
startup_cost += (cpu_operator_cost * numGroupCols) * input_tuples;
|
|
startup_cost += aggcosts->finalCost.startup;
|
|
|
|
total_cost = startup_cost;
|
|
total_cost += aggcosts->finalCost.per_tuple * numGroups;
|
|
/* cost of retrieving from hash table */
|
|
total_cost += cpu_tuple_cost * numGroups;
|
|
output_tuples = numGroups;
|
|
}
|
|
|
|
/*
|
|
* Add the disk costs of hash aggregation that spills to disk.
|
|
*
|
|
* Groups that go into the hash table stay in memory until finalized, so
|
|
* spilling and reprocessing tuples doesn't incur additional invocations
|
|
* of transCost or finalCost. Furthermore, the computed hash value is
|
|
* stored with the spilled tuples, so we don't incur extra invocations of
|
|
* the hash function.
|
|
*
|
|
* Hash Agg begins returning tuples after the first batch is complete.
|
|
* Accrue writes (spilled tuples) to startup_cost and to total_cost;
|
|
* accrue reads only to total_cost.
|
|
*/
|
|
if (aggstrategy == AGG_HASHED || aggstrategy == AGG_MIXED)
|
|
{
|
|
double pages;
|
|
double pages_written = 0.0;
|
|
double pages_read = 0.0;
|
|
double spill_cost;
|
|
double hashentrysize;
|
|
double nbatches;
|
|
Size mem_limit;
|
|
uint64 ngroups_limit;
|
|
int num_partitions;
|
|
int depth;
|
|
|
|
/*
|
|
* Estimate number of batches based on the computed limits. If less
|
|
* than or equal to one, all groups are expected to fit in memory;
|
|
* otherwise we expect to spill.
|
|
*/
|
|
hashentrysize = hash_agg_entry_size(list_length(root->aggtransinfos),
|
|
input_width,
|
|
aggcosts->transitionSpace);
|
|
hash_agg_set_limits(hashentrysize, numGroups, 0, &mem_limit,
|
|
&ngroups_limit, &num_partitions);
|
|
|
|
nbatches = Max((numGroups * hashentrysize) / mem_limit,
|
|
numGroups / ngroups_limit);
|
|
|
|
nbatches = Max(ceil(nbatches), 1.0);
|
|
num_partitions = Max(num_partitions, 2);
|
|
|
|
/*
|
|
* The number of partitions can change at different levels of
|
|
* recursion; but for the purposes of this calculation assume it stays
|
|
* constant.
|
|
*/
|
|
depth = ceil(log(nbatches) / log(num_partitions));
|
|
|
|
/*
|
|
* Estimate number of pages read and written. For each level of
|
|
* recursion, a tuple must be written and then later read.
|
|
*/
|
|
pages = relation_byte_size(input_tuples, input_width) / BLCKSZ;
|
|
pages_written = pages_read = pages * depth;
|
|
|
|
/*
|
|
* HashAgg has somewhat worse IO behavior than Sort on typical
|
|
* hardware/OS combinations. Account for this with a generic penalty.
|
|
*/
|
|
pages_read *= 2.0;
|
|
pages_written *= 2.0;
|
|
|
|
startup_cost += pages_written * random_page_cost;
|
|
total_cost += pages_written * random_page_cost;
|
|
total_cost += pages_read * seq_page_cost;
|
|
|
|
/* account for CPU cost of spilling a tuple and reading it back */
|
|
spill_cost = depth * input_tuples * 2.0 * cpu_tuple_cost;
|
|
startup_cost += spill_cost;
|
|
total_cost += spill_cost;
|
|
}
|
|
|
|
/*
|
|
* If there are quals (HAVING quals), account for their cost and
|
|
* selectivity.
|
|
*/
|
|
if (quals)
|
|
{
|
|
QualCost qual_cost;
|
|
|
|
cost_qual_eval(&qual_cost, quals, root);
|
|
startup_cost += qual_cost.startup;
|
|
total_cost += qual_cost.startup + output_tuples * qual_cost.per_tuple;
|
|
|
|
output_tuples = clamp_row_est(output_tuples *
|
|
clauselist_selectivity(root,
|
|
quals,
|
|
0,
|
|
JOIN_INNER,
|
|
NULL));
|
|
}
|
|
|
|
path->rows = output_tuples;
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = total_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_windowagg
|
|
* Determines and returns the cost of performing a WindowAgg plan node,
|
|
* including the cost of its input.
|
|
*
|
|
* Input is assumed already properly sorted.
|
|
*/
|
|
void
|
|
cost_windowagg(Path *path, PlannerInfo *root,
|
|
List *windowFuncs, int numPartCols, int numOrderCols,
|
|
Cost input_startup_cost, Cost input_total_cost,
|
|
double input_tuples)
|
|
{
|
|
Cost startup_cost;
|
|
Cost total_cost;
|
|
ListCell *lc;
|
|
|
|
startup_cost = input_startup_cost;
|
|
total_cost = input_total_cost;
|
|
|
|
/*
|
|
* Window functions are assumed to cost their stated execution cost, plus
|
|
* the cost of evaluating their input expressions, per tuple. Since they
|
|
* may in fact evaluate their inputs at multiple rows during each cycle,
|
|
* this could be a drastic underestimate; but without a way to know how
|
|
* many rows the window function will fetch, it's hard to do better. In
|
|
* any case, it's a good estimate for all the built-in window functions,
|
|
* so we'll just do this for now.
|
|
*/
|
|
foreach(lc, windowFuncs)
|
|
{
|
|
WindowFunc *wfunc = lfirst_node(WindowFunc, lc);
|
|
Cost wfunccost;
|
|
QualCost argcosts;
|
|
|
|
argcosts.startup = argcosts.per_tuple = 0;
|
|
add_function_cost(root, wfunc->winfnoid, (Node *) wfunc,
|
|
&argcosts);
|
|
startup_cost += argcosts.startup;
|
|
wfunccost = argcosts.per_tuple;
|
|
|
|
/* also add the input expressions' cost to per-input-row costs */
|
|
cost_qual_eval_node(&argcosts, (Node *) wfunc->args, root);
|
|
startup_cost += argcosts.startup;
|
|
wfunccost += argcosts.per_tuple;
|
|
|
|
/*
|
|
* Add the filter's cost to per-input-row costs. XXX We should reduce
|
|
* input expression costs according to filter selectivity.
|
|
*/
|
|
cost_qual_eval_node(&argcosts, (Node *) wfunc->aggfilter, root);
|
|
startup_cost += argcosts.startup;
|
|
wfunccost += argcosts.per_tuple;
|
|
|
|
total_cost += wfunccost * input_tuples;
|
|
}
|
|
|
|
/*
|
|
* We also charge cpu_operator_cost per grouping column per tuple for
|
|
* grouping comparisons, plus cpu_tuple_cost per tuple for general
|
|
* overhead.
|
|
*
|
|
* XXX this neglects costs of spooling the data to disk when it overflows
|
|
* work_mem. Sooner or later that should get accounted for.
|
|
*/
|
|
total_cost += cpu_operator_cost * (numPartCols + numOrderCols) * input_tuples;
|
|
total_cost += cpu_tuple_cost * input_tuples;
|
|
|
|
path->rows = input_tuples;
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = total_cost;
|
|
}
|
|
|
|
/*
|
|
* cost_group
|
|
* Determines and returns the cost of performing a Group plan node,
|
|
* including the cost of its input.
|
|
*
|
|
* Note: caller must ensure that input costs are for appropriately-sorted
|
|
* input.
|
|
*/
|
|
void
|
|
cost_group(Path *path, PlannerInfo *root,
|
|
int numGroupCols, double numGroups,
|
|
List *quals,
|
|
Cost input_startup_cost, Cost input_total_cost,
|
|
double input_tuples)
|
|
{
|
|
double output_tuples;
|
|
Cost startup_cost;
|
|
Cost total_cost;
|
|
|
|
output_tuples = numGroups;
|
|
startup_cost = input_startup_cost;
|
|
total_cost = input_total_cost;
|
|
|
|
/*
|
|
* Charge one cpu_operator_cost per comparison per input tuple. We assume
|
|
* all columns get compared at most of the tuples.
|
|
*/
|
|
total_cost += cpu_operator_cost * input_tuples * numGroupCols;
|
|
|
|
/*
|
|
* If there are quals (HAVING quals), account for their cost and
|
|
* selectivity.
|
|
*/
|
|
if (quals)
|
|
{
|
|
QualCost qual_cost;
|
|
|
|
cost_qual_eval(&qual_cost, quals, root);
|
|
startup_cost += qual_cost.startup;
|
|
total_cost += qual_cost.startup + output_tuples * qual_cost.per_tuple;
|
|
|
|
output_tuples = clamp_row_est(output_tuples *
|
|
clauselist_selectivity(root,
|
|
quals,
|
|
0,
|
|
JOIN_INNER,
|
|
NULL));
|
|
}
|
|
|
|
path->rows = output_tuples;
|
|
path->startup_cost = startup_cost;
|
|
path->total_cost = total_cost;
|
|
}
|
|
|
|
/*
|
|
* initial_cost_nestloop
|
|
* Preliminary estimate of the cost of a nestloop join path.
|
|
*
|
|
* This must quickly produce lower-bound estimates of the path's startup and
|
|
* total costs. If we are unable to eliminate the proposed path from
|
|
* consideration using the lower bounds, final_cost_nestloop will be called
|
|
* to obtain the final estimates.
|
|
*
|
|
* The exact division of labor between this function and final_cost_nestloop
|
|
* is private to them, and represents a tradeoff between speed of the initial
|
|
* estimate and getting a tight lower bound. We choose to not examine the
|
|
* join quals here, since that's by far the most expensive part of the
|
|
* calculations. The end result is that CPU-cost considerations must be
|
|
* left for the second phase; and for SEMI/ANTI joins, we must also postpone
|
|
* incorporation of the inner path's run cost.
|
|
*
|
|
* 'workspace' is to be filled with startup_cost, total_cost, and perhaps
|
|
* other data to be used by final_cost_nestloop
|
|
* 'jointype' is the type of join to be performed
|
|
* 'outer_path' is the outer input to the join
|
|
* 'inner_path' is the inner input to the join
|
|
* 'extra' contains miscellaneous information about the join
|
|
*/
|
|
void
|
|
initial_cost_nestloop(PlannerInfo *root, JoinCostWorkspace *workspace,
|
|
JoinType jointype,
|
|
Path *outer_path, Path *inner_path,
|
|
JoinPathExtraData *extra)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
double outer_path_rows = outer_path->rows;
|
|
Cost inner_rescan_start_cost;
|
|
Cost inner_rescan_total_cost;
|
|
Cost inner_run_cost;
|
|
Cost inner_rescan_run_cost;
|
|
|
|
/* estimate costs to rescan the inner relation */
|
|
cost_rescan(root, inner_path,
|
|
&inner_rescan_start_cost,
|
|
&inner_rescan_total_cost);
|
|
|
|
/* cost of source data */
|
|
|
|
/*
|
|
* NOTE: clearly, we must pay both outer and inner paths' startup_cost
|
|
* before we can start returning tuples, so the join's startup cost is
|
|
* their sum. We'll also pay the inner path's rescan startup cost
|
|
* multiple times.
|
|
*/
|
|
startup_cost += outer_path->startup_cost + inner_path->startup_cost;
|
|
run_cost += outer_path->total_cost - outer_path->startup_cost;
|
|
if (outer_path_rows > 1)
|
|
run_cost += (outer_path_rows - 1) * inner_rescan_start_cost;
|
|
|
|
inner_run_cost = inner_path->total_cost - inner_path->startup_cost;
|
|
inner_rescan_run_cost = inner_rescan_total_cost - inner_rescan_start_cost;
|
|
|
|
if (jointype == JOIN_SEMI || jointype == JOIN_ANTI ||
|
|
extra->inner_unique)
|
|
{
|
|
/*
|
|
* With a SEMI or ANTI join, or if the innerrel is known unique, the
|
|
* executor will stop after the first match.
|
|
*
|
|
* Getting decent estimates requires inspection of the join quals,
|
|
* which we choose to postpone to final_cost_nestloop.
|
|
*/
|
|
|
|
/* Save private data for final_cost_nestloop */
|
|
workspace->inner_run_cost = inner_run_cost;
|
|
workspace->inner_rescan_run_cost = inner_rescan_run_cost;
|
|
}
|
|
else
|
|
{
|
|
/* Normal case; we'll scan whole input rel for each outer row */
|
|
run_cost += inner_run_cost;
|
|
if (outer_path_rows > 1)
|
|
run_cost += (outer_path_rows - 1) * inner_rescan_run_cost;
|
|
}
|
|
|
|
/* CPU costs left for later */
|
|
|
|
/* Public result fields */
|
|
workspace->startup_cost = startup_cost;
|
|
workspace->total_cost = startup_cost + run_cost;
|
|
/* Save private data for final_cost_nestloop */
|
|
workspace->run_cost = run_cost;
|
|
}
|
|
|
|
/*
|
|
* final_cost_nestloop
|
|
* Final estimate of the cost and result size of a nestloop join path.
|
|
*
|
|
* 'path' is already filled in except for the rows and cost fields
|
|
* 'workspace' is the result from initial_cost_nestloop
|
|
* 'extra' contains miscellaneous information about the join
|
|
*/
|
|
void
|
|
final_cost_nestloop(PlannerInfo *root, NestPath *path,
|
|
JoinCostWorkspace *workspace,
|
|
JoinPathExtraData *extra)
|
|
{
|
|
Path *outer_path = path->jpath.outerjoinpath;
|
|
Path *inner_path = path->jpath.innerjoinpath;
|
|
double outer_path_rows = outer_path->rows;
|
|
double inner_path_rows = inner_path->rows;
|
|
Cost startup_cost = workspace->startup_cost;
|
|
Cost run_cost = workspace->run_cost;
|
|
Cost cpu_per_tuple;
|
|
QualCost restrict_qual_cost;
|
|
double ntuples;
|
|
|
|
/* Protect some assumptions below that rowcounts aren't zero */
|
|
if (outer_path_rows <= 0)
|
|
outer_path_rows = 1;
|
|
if (inner_path_rows <= 0)
|
|
inner_path_rows = 1;
|
|
/* Mark the path with the correct row estimate */
|
|
if (path->jpath.path.param_info)
|
|
path->jpath.path.rows = path->jpath.path.param_info->ppi_rows;
|
|
else
|
|
path->jpath.path.rows = path->jpath.path.parent->rows;
|
|
|
|
/* For partial paths, scale row estimate. */
|
|
if (path->jpath.path.parallel_workers > 0)
|
|
{
|
|
double parallel_divisor = get_parallel_divisor(&path->jpath.path);
|
|
|
|
path->jpath.path.rows =
|
|
clamp_row_est(path->jpath.path.rows / parallel_divisor);
|
|
}
|
|
|
|
/*
|
|
* We could include disable_cost in the preliminary estimate, but that
|
|
* would amount to optimizing for the case where the join method is
|
|
* disabled, which doesn't seem like the way to bet.
|
|
*/
|
|
if (!enable_nestloop)
|
|
startup_cost += disable_cost;
|
|
|
|
/* cost of inner-relation source data (we already dealt with outer rel) */
|
|
|
|
if (path->jpath.jointype == JOIN_SEMI || path->jpath.jointype == JOIN_ANTI ||
|
|
extra->inner_unique)
|
|
{
|
|
/*
|
|
* With a SEMI or ANTI join, or if the innerrel is known unique, the
|
|
* executor will stop after the first match.
|
|
*/
|
|
Cost inner_run_cost = workspace->inner_run_cost;
|
|
Cost inner_rescan_run_cost = workspace->inner_rescan_run_cost;
|
|
double outer_matched_rows;
|
|
double outer_unmatched_rows;
|
|
Selectivity inner_scan_frac;
|
|
|
|
/*
|
|
* For an outer-rel row that has at least one match, we can expect the
|
|
* inner scan to stop after a fraction 1/(match_count+1) of the inner
|
|
* rows, if the matches are evenly distributed. Since they probably
|
|
* aren't quite evenly distributed, we apply a fuzz factor of 2.0 to
|
|
* that fraction. (If we used a larger fuzz factor, we'd have to
|
|
* clamp inner_scan_frac to at most 1.0; but since match_count is at
|
|
* least 1, no such clamp is needed now.)
|
|
*/
|
|
outer_matched_rows = rint(outer_path_rows * extra->semifactors.outer_match_frac);
|
|
outer_unmatched_rows = outer_path_rows - outer_matched_rows;
|
|
inner_scan_frac = 2.0 / (extra->semifactors.match_count + 1.0);
|
|
|
|
/*
|
|
* Compute number of tuples processed (not number emitted!). First,
|
|
* account for successfully-matched outer rows.
|
|
*/
|
|
ntuples = outer_matched_rows * inner_path_rows * inner_scan_frac;
|
|
|
|
/*
|
|
* Now we need to estimate the actual costs of scanning the inner
|
|
* relation, which may be quite a bit less than N times inner_run_cost
|
|
* due to early scan stops. We consider two cases. If the inner path
|
|
* is an indexscan using all the joinquals as indexquals, then an
|
|
* unmatched outer row results in an indexscan returning no rows,
|
|
* which is probably quite cheap. Otherwise, the executor will have
|
|
* to scan the whole inner rel for an unmatched row; not so cheap.
|
|
*/
|
|
if (has_indexed_join_quals(path))
|
|
{
|
|
/*
|
|
* Successfully-matched outer rows will only require scanning
|
|
* inner_scan_frac of the inner relation. In this case, we don't
|
|
* need to charge the full inner_run_cost even when that's more
|
|
* than inner_rescan_run_cost, because we can assume that none of
|
|
* the inner scans ever scan the whole inner relation. So it's
|
|
* okay to assume that all the inner scan executions can be
|
|
* fractions of the full cost, even if materialization is reducing
|
|
* the rescan cost. At this writing, it's impossible to get here
|
|
* for a materialized inner scan, so inner_run_cost and
|
|
* inner_rescan_run_cost will be the same anyway; but just in
|
|
* case, use inner_run_cost for the first matched tuple and
|
|
* inner_rescan_run_cost for additional ones.
|
|
*/
|
|
run_cost += inner_run_cost * inner_scan_frac;
|
|
if (outer_matched_rows > 1)
|
|
run_cost += (outer_matched_rows - 1) * inner_rescan_run_cost * inner_scan_frac;
|
|
|
|
/*
|
|
* Add the cost of inner-scan executions for unmatched outer rows.
|
|
* We estimate this as the same cost as returning the first tuple
|
|
* of a nonempty scan. We consider that these are all rescans,
|
|
* since we used inner_run_cost once already.
|
|
*/
|
|
run_cost += outer_unmatched_rows *
|
|
inner_rescan_run_cost / inner_path_rows;
|
|
|
|
/*
|
|
* We won't be evaluating any quals at all for unmatched rows, so
|
|
* don't add them to ntuples.
|
|
*/
|
|
}
|
|
else
|
|
{
|
|
/*
|
|
* Here, a complicating factor is that rescans may be cheaper than
|
|
* first scans. If we never scan all the way to the end of the
|
|
* inner rel, it might be (depending on the plan type) that we'd
|
|
* never pay the whole inner first-scan run cost. However it is
|
|
* difficult to estimate whether that will happen (and it could
|
|
* not happen if there are any unmatched outer rows!), so be
|
|
* conservative and always charge the whole first-scan cost once.
|
|
* We consider this charge to correspond to the first unmatched
|
|
* outer row, unless there isn't one in our estimate, in which
|
|
* case blame it on the first matched row.
|
|
*/
|
|
|
|
/* First, count all unmatched join tuples as being processed */
|
|
ntuples += outer_unmatched_rows * inner_path_rows;
|
|
|
|
/* Now add the forced full scan, and decrement appropriate count */
|
|
run_cost += inner_run_cost;
|
|
if (outer_unmatched_rows >= 1)
|
|
outer_unmatched_rows -= 1;
|
|
else
|
|
outer_matched_rows -= 1;
|
|
|
|
/* Add inner run cost for additional outer tuples having matches */
|
|
if (outer_matched_rows > 0)
|
|
run_cost += outer_matched_rows * inner_rescan_run_cost * inner_scan_frac;
|
|
|
|
/* Add inner run cost for additional unmatched outer tuples */
|
|
if (outer_unmatched_rows > 0)
|
|
run_cost += outer_unmatched_rows * inner_rescan_run_cost;
|
|
}
|
|
}
|
|
else
|
|
{
|
|
/* Normal-case source costs were included in preliminary estimate */
|
|
|
|
/* Compute number of tuples processed (not number emitted!) */
|
|
ntuples = outer_path_rows * inner_path_rows;
|
|
}
|
|
|
|
/* CPU costs */
|
|
cost_qual_eval(&restrict_qual_cost, path->jpath.joinrestrictinfo, root);
|
|
startup_cost += restrict_qual_cost.startup;
|
|
cpu_per_tuple = cpu_tuple_cost + restrict_qual_cost.per_tuple;
|
|
run_cost += cpu_per_tuple * ntuples;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->jpath.path.pathtarget->cost.startup;
|
|
run_cost += path->jpath.path.pathtarget->cost.per_tuple * path->jpath.path.rows;
|
|
|
|
path->jpath.path.startup_cost = startup_cost;
|
|
path->jpath.path.total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* initial_cost_mergejoin
|
|
* Preliminary estimate of the cost of a mergejoin path.
|
|
*
|
|
* This must quickly produce lower-bound estimates of the path's startup and
|
|
* total costs. If we are unable to eliminate the proposed path from
|
|
* consideration using the lower bounds, final_cost_mergejoin will be called
|
|
* to obtain the final estimates.
|
|
*
|
|
* The exact division of labor between this function and final_cost_mergejoin
|
|
* is private to them, and represents a tradeoff between speed of the initial
|
|
* estimate and getting a tight lower bound. We choose to not examine the
|
|
* join quals here, except for obtaining the scan selectivity estimate which
|
|
* is really essential (but fortunately, use of caching keeps the cost of
|
|
* getting that down to something reasonable).
|
|
* We also assume that cost_sort is cheap enough to use here.
|
|
*
|
|
* 'workspace' is to be filled with startup_cost, total_cost, and perhaps
|
|
* other data to be used by final_cost_mergejoin
|
|
* 'jointype' is the type of join to be performed
|
|
* 'mergeclauses' is the list of joinclauses to be used as merge clauses
|
|
* 'outer_path' is the outer input to the join
|
|
* 'inner_path' is the inner input to the join
|
|
* 'outersortkeys' is the list of sort keys for the outer path
|
|
* 'innersortkeys' is the list of sort keys for the inner path
|
|
* 'extra' contains miscellaneous information about the join
|
|
*
|
|
* Note: outersortkeys and innersortkeys should be NIL if no explicit
|
|
* sort is needed because the respective source path is already ordered.
|
|
*/
|
|
void
|
|
initial_cost_mergejoin(PlannerInfo *root, JoinCostWorkspace *workspace,
|
|
JoinType jointype,
|
|
List *mergeclauses,
|
|
Path *outer_path, Path *inner_path,
|
|
List *outersortkeys, List *innersortkeys,
|
|
JoinPathExtraData *extra)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
double outer_path_rows = outer_path->rows;
|
|
double inner_path_rows = inner_path->rows;
|
|
Cost inner_run_cost;
|
|
double outer_rows,
|
|
inner_rows,
|
|
outer_skip_rows,
|
|
inner_skip_rows;
|
|
Selectivity outerstartsel,
|
|
outerendsel,
|
|
innerstartsel,
|
|
innerendsel;
|
|
Path sort_path; /* dummy for result of cost_sort */
|
|
|
|
/* Protect some assumptions below that rowcounts aren't zero */
|
|
if (outer_path_rows <= 0)
|
|
outer_path_rows = 1;
|
|
if (inner_path_rows <= 0)
|
|
inner_path_rows = 1;
|
|
|
|
/*
|
|
* A merge join will stop as soon as it exhausts either input stream
|
|
* (unless it's an outer join, in which case the outer side has to be
|
|
* scanned all the way anyway). Estimate fraction of the left and right
|
|
* inputs that will actually need to be scanned. Likewise, we can
|
|
* estimate the number of rows that will be skipped before the first join
|
|
* pair is found, which should be factored into startup cost. We use only
|
|
* the first (most significant) merge clause for this purpose. Since
|
|
* mergejoinscansel() is a fairly expensive computation, we cache the
|
|
* results in the merge clause RestrictInfo.
|
|
*/
|
|
if (mergeclauses && jointype != JOIN_FULL)
|
|
{
|
|
RestrictInfo *firstclause = (RestrictInfo *) linitial(mergeclauses);
|
|
List *opathkeys;
|
|
List *ipathkeys;
|
|
PathKey *opathkey;
|
|
PathKey *ipathkey;
|
|
MergeScanSelCache *cache;
|
|
|
|
/* Get the input pathkeys to determine the sort-order details */
|
|
opathkeys = outersortkeys ? outersortkeys : outer_path->pathkeys;
|
|
ipathkeys = innersortkeys ? innersortkeys : inner_path->pathkeys;
|
|
Assert(opathkeys);
|
|
Assert(ipathkeys);
|
|
opathkey = (PathKey *) linitial(opathkeys);
|
|
ipathkey = (PathKey *) linitial(ipathkeys);
|
|
/* debugging check */
|
|
if (opathkey->pk_opfamily != ipathkey->pk_opfamily ||
|
|
opathkey->pk_eclass->ec_collation != ipathkey->pk_eclass->ec_collation ||
|
|
opathkey->pk_strategy != ipathkey->pk_strategy ||
|
|
opathkey->pk_nulls_first != ipathkey->pk_nulls_first)
|
|
elog(ERROR, "left and right pathkeys do not match in mergejoin");
|
|
|
|
/* Get the selectivity with caching */
|
|
cache = cached_scansel(root, firstclause, opathkey);
|
|
|
|
if (bms_is_subset(firstclause->left_relids,
|
|
outer_path->parent->relids))
|
|
{
|
|
/* left side of clause is outer */
|
|
outerstartsel = cache->leftstartsel;
|
|
outerendsel = cache->leftendsel;
|
|
innerstartsel = cache->rightstartsel;
|
|
innerendsel = cache->rightendsel;
|
|
}
|
|
else
|
|
{
|
|
/* left side of clause is inner */
|
|
outerstartsel = cache->rightstartsel;
|
|
outerendsel = cache->rightendsel;
|
|
innerstartsel = cache->leftstartsel;
|
|
innerendsel = cache->leftendsel;
|
|
}
|
|
if (jointype == JOIN_LEFT ||
|
|
jointype == JOIN_ANTI)
|
|
{
|
|
outerstartsel = 0.0;
|
|
outerendsel = 1.0;
|
|
}
|
|
else if (jointype == JOIN_RIGHT ||
|
|
jointype == JOIN_RIGHT_ANTI)
|
|
{
|
|
innerstartsel = 0.0;
|
|
innerendsel = 1.0;
|
|
}
|
|
}
|
|
else
|
|
{
|
|
/* cope with clauseless or full mergejoin */
|
|
outerstartsel = innerstartsel = 0.0;
|
|
outerendsel = innerendsel = 1.0;
|
|
}
|
|
|
|
/*
|
|
* Convert selectivities to row counts. We force outer_rows and
|
|
* inner_rows to be at least 1, but the skip_rows estimates can be zero.
|
|
*/
|
|
outer_skip_rows = rint(outer_path_rows * outerstartsel);
|
|
inner_skip_rows = rint(inner_path_rows * innerstartsel);
|
|
outer_rows = clamp_row_est(outer_path_rows * outerendsel);
|
|
inner_rows = clamp_row_est(inner_path_rows * innerendsel);
|
|
|
|
Assert(outer_skip_rows <= outer_rows);
|
|
Assert(inner_skip_rows <= inner_rows);
|
|
|
|
/*
|
|
* Readjust scan selectivities to account for above rounding. This is
|
|
* normally an insignificant effect, but when there are only a few rows in
|
|
* the inputs, failing to do this makes for a large percentage error.
|
|
*/
|
|
outerstartsel = outer_skip_rows / outer_path_rows;
|
|
innerstartsel = inner_skip_rows / inner_path_rows;
|
|
outerendsel = outer_rows / outer_path_rows;
|
|
innerendsel = inner_rows / inner_path_rows;
|
|
|
|
Assert(outerstartsel <= outerendsel);
|
|
Assert(innerstartsel <= innerendsel);
|
|
|
|
/* cost of source data */
|
|
|
|
if (outersortkeys) /* do we need to sort outer? */
|
|
{
|
|
cost_sort(&sort_path,
|
|
root,
|
|
outersortkeys,
|
|
outer_path->total_cost,
|
|
outer_path_rows,
|
|
outer_path->pathtarget->width,
|
|
0.0,
|
|
work_mem,
|
|
-1.0);
|
|
startup_cost += sort_path.startup_cost;
|
|
startup_cost += (sort_path.total_cost - sort_path.startup_cost)
|
|
* outerstartsel;
|
|
run_cost += (sort_path.total_cost - sort_path.startup_cost)
|
|
* (outerendsel - outerstartsel);
|
|
}
|
|
else
|
|
{
|
|
startup_cost += outer_path->startup_cost;
|
|
startup_cost += (outer_path->total_cost - outer_path->startup_cost)
|
|
* outerstartsel;
|
|
run_cost += (outer_path->total_cost - outer_path->startup_cost)
|
|
* (outerendsel - outerstartsel);
|
|
}
|
|
|
|
if (innersortkeys) /* do we need to sort inner? */
|
|
{
|
|
cost_sort(&sort_path,
|
|
root,
|
|
innersortkeys,
|
|
inner_path->total_cost,
|
|
inner_path_rows,
|
|
inner_path->pathtarget->width,
|
|
0.0,
|
|
work_mem,
|
|
-1.0);
|
|
startup_cost += sort_path.startup_cost;
|
|
startup_cost += (sort_path.total_cost - sort_path.startup_cost)
|
|
* innerstartsel;
|
|
inner_run_cost = (sort_path.total_cost - sort_path.startup_cost)
|
|
* (innerendsel - innerstartsel);
|
|
}
|
|
else
|
|
{
|
|
startup_cost += inner_path->startup_cost;
|
|
startup_cost += (inner_path->total_cost - inner_path->startup_cost)
|
|
* innerstartsel;
|
|
inner_run_cost = (inner_path->total_cost - inner_path->startup_cost)
|
|
* (innerendsel - innerstartsel);
|
|
}
|
|
|
|
/*
|
|
* We can't yet determine whether rescanning occurs, or whether
|
|
* materialization of the inner input should be done. The minimum
|
|
* possible inner input cost, regardless of rescan and materialization
|
|
* considerations, is inner_run_cost. We include that in
|
|
* workspace->total_cost, but not yet in run_cost.
|
|
*/
|
|
|
|
/* CPU costs left for later */
|
|
|
|
/* Public result fields */
|
|
workspace->startup_cost = startup_cost;
|
|
workspace->total_cost = startup_cost + run_cost + inner_run_cost;
|
|
/* Save private data for final_cost_mergejoin */
|
|
workspace->run_cost = run_cost;
|
|
workspace->inner_run_cost = inner_run_cost;
|
|
workspace->outer_rows = outer_rows;
|
|
workspace->inner_rows = inner_rows;
|
|
workspace->outer_skip_rows = outer_skip_rows;
|
|
workspace->inner_skip_rows = inner_skip_rows;
|
|
}
|
|
|
|
/*
|
|
* final_cost_mergejoin
|
|
* Final estimate of the cost and result size of a mergejoin path.
|
|
*
|
|
* Unlike other costsize functions, this routine makes two actual decisions:
|
|
* whether the executor will need to do mark/restore, and whether we should
|
|
* materialize the inner path. It would be logically cleaner to build
|
|
* separate paths testing these alternatives, but that would require repeating
|
|
* most of the cost calculations, which are not all that cheap. Since the
|
|
* choice will not affect output pathkeys or startup cost, only total cost,
|
|
* there is no possibility of wanting to keep more than one path. So it seems
|
|
* best to make the decisions here and record them in the path's
|
|
* skip_mark_restore and materialize_inner fields.
|
|
*
|
|
* Mark/restore overhead is usually required, but can be skipped if we know
|
|
* that the executor need find only one match per outer tuple, and that the
|
|
* mergeclauses are sufficient to identify a match.
|
|
*
|
|
* We materialize the inner path if we need mark/restore and either the inner
|
|
* path can't support mark/restore, or it's cheaper to use an interposed
|
|
* Material node to handle mark/restore.
|
|
*
|
|
* 'path' is already filled in except for the rows and cost fields and
|
|
* skip_mark_restore and materialize_inner
|
|
* 'workspace' is the result from initial_cost_mergejoin
|
|
* 'extra' contains miscellaneous information about the join
|
|
*/
|
|
void
|
|
final_cost_mergejoin(PlannerInfo *root, MergePath *path,
|
|
JoinCostWorkspace *workspace,
|
|
JoinPathExtraData *extra)
|
|
{
|
|
Path *outer_path = path->jpath.outerjoinpath;
|
|
Path *inner_path = path->jpath.innerjoinpath;
|
|
double inner_path_rows = inner_path->rows;
|
|
List *mergeclauses = path->path_mergeclauses;
|
|
List *innersortkeys = path->innersortkeys;
|
|
Cost startup_cost = workspace->startup_cost;
|
|
Cost run_cost = workspace->run_cost;
|
|
Cost inner_run_cost = workspace->inner_run_cost;
|
|
double outer_rows = workspace->outer_rows;
|
|
double inner_rows = workspace->inner_rows;
|
|
double outer_skip_rows = workspace->outer_skip_rows;
|
|
double inner_skip_rows = workspace->inner_skip_rows;
|
|
Cost cpu_per_tuple,
|
|
bare_inner_cost,
|
|
mat_inner_cost;
|
|
QualCost merge_qual_cost;
|
|
QualCost qp_qual_cost;
|
|
double mergejointuples,
|
|
rescannedtuples;
|
|
double rescanratio;
|
|
|
|
/* Protect some assumptions below that rowcounts aren't zero */
|
|
if (inner_path_rows <= 0)
|
|
inner_path_rows = 1;
|
|
|
|
/* Mark the path with the correct row estimate */
|
|
if (path->jpath.path.param_info)
|
|
path->jpath.path.rows = path->jpath.path.param_info->ppi_rows;
|
|
else
|
|
path->jpath.path.rows = path->jpath.path.parent->rows;
|
|
|
|
/* For partial paths, scale row estimate. */
|
|
if (path->jpath.path.parallel_workers > 0)
|
|
{
|
|
double parallel_divisor = get_parallel_divisor(&path->jpath.path);
|
|
|
|
path->jpath.path.rows =
|
|
clamp_row_est(path->jpath.path.rows / parallel_divisor);
|
|
}
|
|
|
|
/*
|
|
* We could include disable_cost in the preliminary estimate, but that
|
|
* would amount to optimizing for the case where the join method is
|
|
* disabled, which doesn't seem like the way to bet.
|
|
*/
|
|
if (!enable_mergejoin)
|
|
startup_cost += disable_cost;
|
|
|
|
/*
|
|
* Compute cost of the mergequals and qpquals (other restriction clauses)
|
|
* separately.
|
|
*/
|
|
cost_qual_eval(&merge_qual_cost, mergeclauses, root);
|
|
cost_qual_eval(&qp_qual_cost, path->jpath.joinrestrictinfo, root);
|
|
qp_qual_cost.startup -= merge_qual_cost.startup;
|
|
qp_qual_cost.per_tuple -= merge_qual_cost.per_tuple;
|
|
|
|
/*
|
|
* With a SEMI or ANTI join, or if the innerrel is known unique, the
|
|
* executor will stop scanning for matches after the first match. When
|
|
* all the joinclauses are merge clauses, this means we don't ever need to
|
|
* back up the merge, and so we can skip mark/restore overhead.
|
|
*/
|
|
if ((path->jpath.jointype == JOIN_SEMI ||
|
|
path->jpath.jointype == JOIN_ANTI ||
|
|
extra->inner_unique) &&
|
|
(list_length(path->jpath.joinrestrictinfo) ==
|
|
list_length(path->path_mergeclauses)))
|
|
path->skip_mark_restore = true;
|
|
else
|
|
path->skip_mark_restore = false;
|
|
|
|
/*
|
|
* Get approx # tuples passing the mergequals. We use approx_tuple_count
|
|
* here because we need an estimate done with JOIN_INNER semantics.
|
|
*/
|
|
mergejointuples = approx_tuple_count(root, &path->jpath, mergeclauses);
|
|
|
|
/*
|
|
* When there are equal merge keys in the outer relation, the mergejoin
|
|
* must rescan any matching tuples in the inner relation. This means
|
|
* re-fetching inner tuples; we have to estimate how often that happens.
|
|
*
|
|
* For regular inner and outer joins, the number of re-fetches can be
|
|
* estimated approximately as size of merge join output minus size of
|
|
* inner relation. Assume that the distinct key values are 1, 2, ..., and
|
|
* denote the number of values of each key in the outer relation as m1,
|
|
* m2, ...; in the inner relation, n1, n2, ... Then we have
|
|
*
|
|
* size of join = m1 * n1 + m2 * n2 + ...
|
|
*
|
|
* number of rescanned tuples = (m1 - 1) * n1 + (m2 - 1) * n2 + ... = m1 *
|
|
* n1 + m2 * n2 + ... - (n1 + n2 + ...) = size of join - size of inner
|
|
* relation
|
|
*
|
|
* This equation works correctly for outer tuples having no inner match
|
|
* (nk = 0), but not for inner tuples having no outer match (mk = 0); we
|
|
* are effectively subtracting those from the number of rescanned tuples,
|
|
* when we should not. Can we do better without expensive selectivity
|
|
* computations?
|
|
*
|
|
* The whole issue is moot if we are working from a unique-ified outer
|
|
* input, or if we know we don't need to mark/restore at all.
|
|
*/
|
|
if (IsA(outer_path, UniquePath) || path->skip_mark_restore)
|
|
rescannedtuples = 0;
|
|
else
|
|
{
|
|
rescannedtuples = mergejointuples - inner_path_rows;
|
|
/* Must clamp because of possible underestimate */
|
|
if (rescannedtuples < 0)
|
|
rescannedtuples = 0;
|
|
}
|
|
|
|
/*
|
|
* We'll inflate various costs this much to account for rescanning. Note
|
|
* that this is to be multiplied by something involving inner_rows, or
|
|
* another number related to the portion of the inner rel we'll scan.
|
|
*/
|
|
rescanratio = 1.0 + (rescannedtuples / inner_rows);
|
|
|
|
/*
|
|
* Decide whether we want to materialize the inner input to shield it from
|
|
* mark/restore and performing re-fetches. Our cost model for regular
|
|
* re-fetches is that a re-fetch costs the same as an original fetch,
|
|
* which is probably an overestimate; but on the other hand we ignore the
|
|
* bookkeeping costs of mark/restore. Not clear if it's worth developing
|
|
* a more refined model. So we just need to inflate the inner run cost by
|
|
* rescanratio.
|
|
*/
|
|
bare_inner_cost = inner_run_cost * rescanratio;
|
|
|
|
/*
|
|
* When we interpose a Material node the re-fetch cost is assumed to be
|
|
* just cpu_operator_cost per tuple, independently of the underlying
|
|
* plan's cost; and we charge an extra cpu_operator_cost per original
|
|
* fetch as well. Note that we're assuming the materialize node will
|
|
* never spill to disk, since it only has to remember tuples back to the
|
|
* last mark. (If there are a huge number of duplicates, our other cost
|
|
* factors will make the path so expensive that it probably won't get
|
|
* chosen anyway.) So we don't use cost_rescan here.
|
|
*
|
|
* Note: keep this estimate in sync with create_mergejoin_plan's labeling
|
|
* of the generated Material node.
|
|
*/
|
|
mat_inner_cost = inner_run_cost +
|
|
cpu_operator_cost * inner_rows * rescanratio;
|
|
|
|
/*
|
|
* If we don't need mark/restore at all, we don't need materialization.
|
|
*/
|
|
if (path->skip_mark_restore)
|
|
path->materialize_inner = false;
|
|
|
|
/*
|
|
* Prefer materializing if it looks cheaper, unless the user has asked to
|
|
* suppress materialization.
|
|
*/
|
|
else if (enable_material && mat_inner_cost < bare_inner_cost)
|
|
path->materialize_inner = true;
|
|
|
|
/*
|
|
* Even if materializing doesn't look cheaper, we *must* do it if the
|
|
* inner path is to be used directly (without sorting) and it doesn't
|
|
* support mark/restore.
|
|
*
|
|
* Since the inner side must be ordered, and only Sorts and IndexScans can
|
|
* create order to begin with, and they both support mark/restore, you
|
|
* might think there's no problem --- but you'd be wrong. Nestloop and
|
|
* merge joins can *preserve* the order of their inputs, so they can be
|
|
* selected as the input of a mergejoin, and they don't support
|
|
* mark/restore at present.
|
|
*
|
|
* We don't test the value of enable_material here, because
|
|
* materialization is required for correctness in this case, and turning
|
|
* it off does not entitle us to deliver an invalid plan.
|
|
*/
|
|
else if (innersortkeys == NIL &&
|
|
!ExecSupportsMarkRestore(inner_path))
|
|
path->materialize_inner = true;
|
|
|
|
/*
|
|
* Also, force materializing if the inner path is to be sorted and the
|
|
* sort is expected to spill to disk. This is because the final merge
|
|
* pass can be done on-the-fly if it doesn't have to support mark/restore.
|
|
* We don't try to adjust the cost estimates for this consideration,
|
|
* though.
|
|
*
|
|
* Since materialization is a performance optimization in this case,
|
|
* rather than necessary for correctness, we skip it if enable_material is
|
|
* off.
|
|
*/
|
|
else if (enable_material && innersortkeys != NIL &&
|
|
relation_byte_size(inner_path_rows,
|
|
inner_path->pathtarget->width) >
|
|
(work_mem * 1024L))
|
|
path->materialize_inner = true;
|
|
else
|
|
path->materialize_inner = false;
|
|
|
|
/* Charge the right incremental cost for the chosen case */
|
|
if (path->materialize_inner)
|
|
run_cost += mat_inner_cost;
|
|
else
|
|
run_cost += bare_inner_cost;
|
|
|
|
/* CPU costs */
|
|
|
|
/*
|
|
* The number of tuple comparisons needed is approximately number of outer
|
|
* rows plus number of inner rows plus number of rescanned tuples (can we
|
|
* refine this?). At each one, we need to evaluate the mergejoin quals.
|
|
*/
|
|
startup_cost += merge_qual_cost.startup;
|
|
startup_cost += merge_qual_cost.per_tuple *
|
|
(outer_skip_rows + inner_skip_rows * rescanratio);
|
|
run_cost += merge_qual_cost.per_tuple *
|
|
((outer_rows - outer_skip_rows) +
|
|
(inner_rows - inner_skip_rows) * rescanratio);
|
|
|
|
/*
|
|
* For each tuple that gets through the mergejoin proper, we charge
|
|
* cpu_tuple_cost plus the cost of evaluating additional restriction
|
|
* clauses that are to be applied at the join. (This is pessimistic since
|
|
* not all of the quals may get evaluated at each tuple.)
|
|
*
|
|
* Note: we could adjust for SEMI/ANTI joins skipping some qual
|
|
* evaluations here, but it's probably not worth the trouble.
|
|
*/
|
|
startup_cost += qp_qual_cost.startup;
|
|
cpu_per_tuple = cpu_tuple_cost + qp_qual_cost.per_tuple;
|
|
run_cost += cpu_per_tuple * mergejointuples;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->jpath.path.pathtarget->cost.startup;
|
|
run_cost += path->jpath.path.pathtarget->cost.per_tuple * path->jpath.path.rows;
|
|
|
|
path->jpath.path.startup_cost = startup_cost;
|
|
path->jpath.path.total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
/*
|
|
* run mergejoinscansel() with caching
|
|
*/
|
|
static MergeScanSelCache *
|
|
cached_scansel(PlannerInfo *root, RestrictInfo *rinfo, PathKey *pathkey)
|
|
{
|
|
MergeScanSelCache *cache;
|
|
ListCell *lc;
|
|
Selectivity leftstartsel,
|
|
leftendsel,
|
|
rightstartsel,
|
|
rightendsel;
|
|
MemoryContext oldcontext;
|
|
|
|
/* Do we have this result already? */
|
|
foreach(lc, rinfo->scansel_cache)
|
|
{
|
|
cache = (MergeScanSelCache *) lfirst(lc);
|
|
if (cache->opfamily == pathkey->pk_opfamily &&
|
|
cache->collation == pathkey->pk_eclass->ec_collation &&
|
|
cache->strategy == pathkey->pk_strategy &&
|
|
cache->nulls_first == pathkey->pk_nulls_first)
|
|
return cache;
|
|
}
|
|
|
|
/* Nope, do the computation */
|
|
mergejoinscansel(root,
|
|
(Node *) rinfo->clause,
|
|
pathkey->pk_opfamily,
|
|
pathkey->pk_strategy,
|
|
pathkey->pk_nulls_first,
|
|
&leftstartsel,
|
|
&leftendsel,
|
|
&rightstartsel,
|
|
&rightendsel);
|
|
|
|
/* Cache the result in suitably long-lived workspace */
|
|
oldcontext = MemoryContextSwitchTo(root->planner_cxt);
|
|
|
|
cache = (MergeScanSelCache *) palloc(sizeof(MergeScanSelCache));
|
|
cache->opfamily = pathkey->pk_opfamily;
|
|
cache->collation = pathkey->pk_eclass->ec_collation;
|
|
cache->strategy = pathkey->pk_strategy;
|
|
cache->nulls_first = pathkey->pk_nulls_first;
|
|
cache->leftstartsel = leftstartsel;
|
|
cache->leftendsel = leftendsel;
|
|
cache->rightstartsel = rightstartsel;
|
|
cache->rightendsel = rightendsel;
|
|
|
|
rinfo->scansel_cache = lappend(rinfo->scansel_cache, cache);
|
|
|
|
MemoryContextSwitchTo(oldcontext);
|
|
|
|
return cache;
|
|
}
|
|
|
|
/*
|
|
* initial_cost_hashjoin
|
|
* Preliminary estimate of the cost of a hashjoin path.
|
|
*
|
|
* This must quickly produce lower-bound estimates of the path's startup and
|
|
* total costs. If we are unable to eliminate the proposed path from
|
|
* consideration using the lower bounds, final_cost_hashjoin will be called
|
|
* to obtain the final estimates.
|
|
*
|
|
* The exact division of labor between this function and final_cost_hashjoin
|
|
* is private to them, and represents a tradeoff between speed of the initial
|
|
* estimate and getting a tight lower bound. We choose to not examine the
|
|
* join quals here (other than by counting the number of hash clauses),
|
|
* so we can't do much with CPU costs. We do assume that
|
|
* ExecChooseHashTableSize is cheap enough to use here.
|
|
*
|
|
* 'workspace' is to be filled with startup_cost, total_cost, and perhaps
|
|
* other data to be used by final_cost_hashjoin
|
|
* 'jointype' is the type of join to be performed
|
|
* 'hashclauses' is the list of joinclauses to be used as hash clauses
|
|
* 'outer_path' is the outer input to the join
|
|
* 'inner_path' is the inner input to the join
|
|
* 'extra' contains miscellaneous information about the join
|
|
* 'parallel_hash' indicates that inner_path is partial and that a shared
|
|
* hash table will be built in parallel
|
|
*/
|
|
void
|
|
initial_cost_hashjoin(PlannerInfo *root, JoinCostWorkspace *workspace,
|
|
JoinType jointype,
|
|
List *hashclauses,
|
|
Path *outer_path, Path *inner_path,
|
|
JoinPathExtraData *extra,
|
|
bool parallel_hash)
|
|
{
|
|
Cost startup_cost = 0;
|
|
Cost run_cost = 0;
|
|
double outer_path_rows = outer_path->rows;
|
|
double inner_path_rows = inner_path->rows;
|
|
double inner_path_rows_total = inner_path_rows;
|
|
int num_hashclauses = list_length(hashclauses);
|
|
int numbuckets;
|
|
int numbatches;
|
|
int num_skew_mcvs;
|
|
size_t space_allowed; /* unused */
|
|
|
|
/* cost of source data */
|
|
startup_cost += outer_path->startup_cost;
|
|
run_cost += outer_path->total_cost - outer_path->startup_cost;
|
|
startup_cost += inner_path->total_cost;
|
|
|
|
/*
|
|
* Cost of computing hash function: must do it once per input tuple. We
|
|
* charge one cpu_operator_cost for each column's hash function. Also,
|
|
* tack on one cpu_tuple_cost per inner row, to model the costs of
|
|
* inserting the row into the hashtable.
|
|
*
|
|
* XXX when a hashclause is more complex than a single operator, we really
|
|
* should charge the extra eval costs of the left or right side, as
|
|
* appropriate, here. This seems more work than it's worth at the moment.
|
|
*/
|
|
startup_cost += (cpu_operator_cost * num_hashclauses + cpu_tuple_cost)
|
|
* inner_path_rows;
|
|
run_cost += cpu_operator_cost * num_hashclauses * outer_path_rows;
|
|
|
|
/*
|
|
* If this is a parallel hash build, then the value we have for
|
|
* inner_rows_total currently refers only to the rows returned by each
|
|
* participant. For shared hash table size estimation, we need the total
|
|
* number, so we need to undo the division.
|
|
*/
|
|
if (parallel_hash)
|
|
inner_path_rows_total *= get_parallel_divisor(inner_path);
|
|
|
|
/*
|
|
* Get hash table size that executor would use for inner relation.
|
|
*
|
|
* XXX for the moment, always assume that skew optimization will be
|
|
* performed. As long as SKEW_HASH_MEM_PERCENT is small, it's not worth
|
|
* trying to determine that for sure.
|
|
*
|
|
* XXX at some point it might be interesting to try to account for skew
|
|
* optimization in the cost estimate, but for now, we don't.
|
|
*/
|
|
ExecChooseHashTableSize(inner_path_rows_total,
|
|
inner_path->pathtarget->width,
|
|
true, /* useskew */
|
|
parallel_hash, /* try_combined_hash_mem */
|
|
outer_path->parallel_workers,
|
|
&space_allowed,
|
|
&numbuckets,
|
|
&numbatches,
|
|
&num_skew_mcvs);
|
|
|
|
/*
|
|
* If inner relation is too big then we will need to "batch" the join,
|
|
* which implies writing and reading most of the tuples to disk an extra
|
|
* time. Charge seq_page_cost per page, since the I/O should be nice and
|
|
* sequential. Writing the inner rel counts as startup cost, all the rest
|
|
* as run cost.
|
|
*/
|
|
if (numbatches > 1)
|
|
{
|
|
double outerpages = page_size(outer_path_rows,
|
|
outer_path->pathtarget->width);
|
|
double innerpages = page_size(inner_path_rows,
|
|
inner_path->pathtarget->width);
|
|
|
|
startup_cost += seq_page_cost * innerpages;
|
|
run_cost += seq_page_cost * (innerpages + 2 * outerpages);
|
|
}
|
|
|
|
/* CPU costs left for later */
|
|
|
|
/* Public result fields */
|
|
workspace->startup_cost = startup_cost;
|
|
workspace->total_cost = startup_cost + run_cost;
|
|
/* Save private data for final_cost_hashjoin */
|
|
workspace->run_cost = run_cost;
|
|
workspace->numbuckets = numbuckets;
|
|
workspace->numbatches = numbatches;
|
|
workspace->inner_rows_total = inner_path_rows_total;
|
|
}
|
|
|
|
/*
|
|
* final_cost_hashjoin
|
|
* Final estimate of the cost and result size of a hashjoin path.
|
|
*
|
|
* Note: the numbatches estimate is also saved into 'path' for use later
|
|
*
|
|
* 'path' is already filled in except for the rows and cost fields and
|
|
* num_batches
|
|
* 'workspace' is the result from initial_cost_hashjoin
|
|
* 'extra' contains miscellaneous information about the join
|
|
*/
|
|
void
|
|
final_cost_hashjoin(PlannerInfo *root, HashPath *path,
|
|
JoinCostWorkspace *workspace,
|
|
JoinPathExtraData *extra)
|
|
{
|
|
Path *outer_path = path->jpath.outerjoinpath;
|
|
Path *inner_path = path->jpath.innerjoinpath;
|
|
double outer_path_rows = outer_path->rows;
|
|
double inner_path_rows = inner_path->rows;
|
|
double inner_path_rows_total = workspace->inner_rows_total;
|
|
List *hashclauses = path->path_hashclauses;
|
|
Cost startup_cost = workspace->startup_cost;
|
|
Cost run_cost = workspace->run_cost;
|
|
int numbuckets = workspace->numbuckets;
|
|
int numbatches = workspace->numbatches;
|
|
Cost cpu_per_tuple;
|
|
QualCost hash_qual_cost;
|
|
QualCost qp_qual_cost;
|
|
double hashjointuples;
|
|
double virtualbuckets;
|
|
Selectivity innerbucketsize;
|
|
Selectivity innermcvfreq;
|
|
ListCell *hcl;
|
|
|
|
/* Mark the path with the correct row estimate */
|
|
if (path->jpath.path.param_info)
|
|
path->jpath.path.rows = path->jpath.path.param_info->ppi_rows;
|
|
else
|
|
path->jpath.path.rows = path->jpath.path.parent->rows;
|
|
|
|
/* For partial paths, scale row estimate. */
|
|
if (path->jpath.path.parallel_workers > 0)
|
|
{
|
|
double parallel_divisor = get_parallel_divisor(&path->jpath.path);
|
|
|
|
path->jpath.path.rows =
|
|
clamp_row_est(path->jpath.path.rows / parallel_divisor);
|
|
}
|
|
|
|
/*
|
|
* We could include disable_cost in the preliminary estimate, but that
|
|
* would amount to optimizing for the case where the join method is
|
|
* disabled, which doesn't seem like the way to bet.
|
|
*/
|
|
if (!enable_hashjoin)
|
|
startup_cost += disable_cost;
|
|
|
|
/* mark the path with estimated # of batches */
|
|
path->num_batches = numbatches;
|
|
|
|
/* store the total number of tuples (sum of partial row estimates) */
|
|
path->inner_rows_total = inner_path_rows_total;
|
|
|
|
/* and compute the number of "virtual" buckets in the whole join */
|
|
virtualbuckets = (double) numbuckets * (double) numbatches;
|
|
|
|
/*
|
|
* Determine bucketsize fraction and MCV frequency for the inner relation.
|
|
* We use the smallest bucketsize or MCV frequency estimated for any
|
|
* individual hashclause; this is undoubtedly conservative.
|
|
*
|
|
* BUT: if inner relation has been unique-ified, we can assume it's good
|
|
* for hashing. This is important both because it's the right answer, and
|
|
* because we avoid contaminating the cache with a value that's wrong for
|
|
* non-unique-ified paths.
|
|
*/
|
|
if (IsA(inner_path, UniquePath))
|
|
{
|
|
innerbucketsize = 1.0 / virtualbuckets;
|
|
innermcvfreq = 0.0;
|
|
}
|
|
else
|
|
{
|
|
innerbucketsize = 1.0;
|
|
innermcvfreq = 1.0;
|
|
foreach(hcl, hashclauses)
|
|
{
|
|
RestrictInfo *restrictinfo = lfirst_node(RestrictInfo, hcl);
|
|
Selectivity thisbucketsize;
|
|
Selectivity thismcvfreq;
|
|
|
|
/*
|
|
* First we have to figure out which side of the hashjoin clause
|
|
* is the inner side.
|
|
*
|
|
* Since we tend to visit the same clauses over and over when
|
|
* planning a large query, we cache the bucket stats estimates in
|
|
* the RestrictInfo node to avoid repeated lookups of statistics.
|
|
*/
|
|
if (bms_is_subset(restrictinfo->right_relids,
|
|
inner_path->parent->relids))
|
|
{
|
|
/* righthand side is inner */
|
|
thisbucketsize = restrictinfo->right_bucketsize;
|
|
if (thisbucketsize < 0)
|
|
{
|
|
/* not cached yet */
|
|
estimate_hash_bucket_stats(root,
|
|
get_rightop(restrictinfo->clause),
|
|
virtualbuckets,
|
|
&restrictinfo->right_mcvfreq,
|
|
&restrictinfo->right_bucketsize);
|
|
thisbucketsize = restrictinfo->right_bucketsize;
|
|
}
|
|
thismcvfreq = restrictinfo->right_mcvfreq;
|
|
}
|
|
else
|
|
{
|
|
Assert(bms_is_subset(restrictinfo->left_relids,
|
|
inner_path->parent->relids));
|
|
/* lefthand side is inner */
|
|
thisbucketsize = restrictinfo->left_bucketsize;
|
|
if (thisbucketsize < 0)
|
|
{
|
|
/* not cached yet */
|
|
estimate_hash_bucket_stats(root,
|
|
get_leftop(restrictinfo->clause),
|
|
virtualbuckets,
|
|
&restrictinfo->left_mcvfreq,
|
|
&restrictinfo->left_bucketsize);
|
|
thisbucketsize = restrictinfo->left_bucketsize;
|
|
}
|
|
thismcvfreq = restrictinfo->left_mcvfreq;
|
|
}
|
|
|
|
if (innerbucketsize > thisbucketsize)
|
|
innerbucketsize = thisbucketsize;
|
|
if (innermcvfreq > thismcvfreq)
|
|
innermcvfreq = thismcvfreq;
|
|
}
|
|
}
|
|
|
|
/*
|
|
* If the bucket holding the inner MCV would exceed hash_mem, we don't
|
|
* want to hash unless there is really no other alternative, so apply
|
|
* disable_cost. (The executor normally copes with excessive memory usage
|
|
* by splitting batches, but obviously it cannot separate equal values
|
|
* that way, so it will be unable to drive the batch size below hash_mem
|
|
* when this is true.)
|
|
*/
|
|
if (relation_byte_size(clamp_row_est(inner_path_rows * innermcvfreq),
|
|
inner_path->pathtarget->width) > get_hash_memory_limit())
|
|
startup_cost += disable_cost;
|
|
|
|
/*
|
|
* Compute cost of the hashquals and qpquals (other restriction clauses)
|
|
* separately.
|
|
*/
|
|
cost_qual_eval(&hash_qual_cost, hashclauses, root);
|
|
cost_qual_eval(&qp_qual_cost, path->jpath.joinrestrictinfo, root);
|
|
qp_qual_cost.startup -= hash_qual_cost.startup;
|
|
qp_qual_cost.per_tuple -= hash_qual_cost.per_tuple;
|
|
|
|
/* CPU costs */
|
|
|
|
if (path->jpath.jointype == JOIN_SEMI ||
|
|
path->jpath.jointype == JOIN_ANTI ||
|
|
extra->inner_unique)
|
|
{
|
|
double outer_matched_rows;
|
|
Selectivity inner_scan_frac;
|
|
|
|
/*
|
|
* With a SEMI or ANTI join, or if the innerrel is known unique, the
|
|
* executor will stop after the first match.
|
|
*
|
|
* For an outer-rel row that has at least one match, we can expect the
|
|
* bucket scan to stop after a fraction 1/(match_count+1) of the
|
|
* bucket's rows, if the matches are evenly distributed. Since they
|
|
* probably aren't quite evenly distributed, we apply a fuzz factor of
|
|
* 2.0 to that fraction. (If we used a larger fuzz factor, we'd have
|
|
* to clamp inner_scan_frac to at most 1.0; but since match_count is
|
|
* at least 1, no such clamp is needed now.)
|
|
*/
|
|
outer_matched_rows = rint(outer_path_rows * extra->semifactors.outer_match_frac);
|
|
inner_scan_frac = 2.0 / (extra->semifactors.match_count + 1.0);
|
|
|
|
startup_cost += hash_qual_cost.startup;
|
|
run_cost += hash_qual_cost.per_tuple * outer_matched_rows *
|
|
clamp_row_est(inner_path_rows * innerbucketsize * inner_scan_frac) * 0.5;
|
|
|
|
/*
|
|
* For unmatched outer-rel rows, the picture is quite a lot different.
|
|
* In the first place, there is no reason to assume that these rows
|
|
* preferentially hit heavily-populated buckets; instead assume they
|
|
* are uncorrelated with the inner distribution and so they see an
|
|
* average bucket size of inner_path_rows / virtualbuckets. In the
|
|
* second place, it seems likely that they will have few if any exact
|
|
* hash-code matches and so very few of the tuples in the bucket will
|
|
* actually require eval of the hash quals. We don't have any good
|
|
* way to estimate how many will, but for the moment assume that the
|
|
* effective cost per bucket entry is one-tenth what it is for
|
|
* matchable tuples.
|
|
*/
|
|
run_cost += hash_qual_cost.per_tuple *
|
|
(outer_path_rows - outer_matched_rows) *
|
|
clamp_row_est(inner_path_rows / virtualbuckets) * 0.05;
|
|
|
|
/* Get # of tuples that will pass the basic join */
|
|
if (path->jpath.jointype == JOIN_ANTI)
|
|
hashjointuples = outer_path_rows - outer_matched_rows;
|
|
else
|
|
hashjointuples = outer_matched_rows;
|
|
}
|
|
else
|
|
{
|
|
/*
|
|
* The number of tuple comparisons needed is the number of outer
|
|
* tuples times the typical number of tuples in a hash bucket, which
|
|
* is the inner relation size times its bucketsize fraction. At each
|
|
* one, we need to evaluate the hashjoin quals. But actually,
|
|
* charging the full qual eval cost at each tuple is pessimistic,
|
|
* since we don't evaluate the quals unless the hash values match
|
|
* exactly. For lack of a better idea, halve the cost estimate to
|
|
* allow for that.
|
|
*/
|
|
startup_cost += hash_qual_cost.startup;
|
|
run_cost += hash_qual_cost.per_tuple * outer_path_rows *
|
|
clamp_row_est(inner_path_rows * innerbucketsize) * 0.5;
|
|
|
|
/*
|
|
* Get approx # tuples passing the hashquals. We use
|
|
* approx_tuple_count here because we need an estimate done with
|
|
* JOIN_INNER semantics.
|
|
*/
|
|
hashjointuples = approx_tuple_count(root, &path->jpath, hashclauses);
|
|
}
|
|
|
|
/*
|
|
* For each tuple that gets through the hashjoin proper, we charge
|
|
* cpu_tuple_cost plus the cost of evaluating additional restriction
|
|
* clauses that are to be applied at the join. (This is pessimistic since
|
|
* not all of the quals may get evaluated at each tuple.)
|
|
*/
|
|
startup_cost += qp_qual_cost.startup;
|
|
cpu_per_tuple = cpu_tuple_cost + qp_qual_cost.per_tuple;
|
|
run_cost += cpu_per_tuple * hashjointuples;
|
|
|
|
/* tlist eval costs are paid per output row, not per tuple scanned */
|
|
startup_cost += path->jpath.path.pathtarget->cost.startup;
|
|
run_cost += path->jpath.path.pathtarget->cost.per_tuple * path->jpath.path.rows;
|
|
|
|
path->jpath.path.startup_cost = startup_cost;
|
|
path->jpath.path.total_cost = startup_cost + run_cost;
|
|
}
|
|
|
|
|
|
/*
|
|
* cost_subplan
|
|
* Figure the costs for a SubPlan (or initplan).
|
|
*
|
|
* Note: we could dig the subplan's Plan out of the root list, but in practice
|
|
* all callers have it handy already, so we make them pass it.
|
|
*/
|
|
void
|
|
cost_subplan(PlannerInfo *root, SubPlan *subplan, Plan *plan)
|
|
{
|
|
QualCost sp_cost;
|
|
|
|
/* Figure any cost for evaluating the testexpr */
|
|
cost_qual_eval(&sp_cost,
|
|
make_ands_implicit((Expr *) subplan->testexpr),
|
|
root);
|
|
|
|
if (subplan->useHashTable)
|
|
{
|
|
/*
|
|
* If we are using a hash table for the subquery outputs, then the
|
|
* cost of evaluating the query is a one-time cost. We charge one
|
|
* cpu_operator_cost per tuple for the work of loading the hashtable,
|
|
* too.
|
|
*/
|
|
sp_cost.startup += plan->total_cost +
|
|
cpu_operator_cost * plan->plan_rows;
|
|
|
|
/*
|
|
* The per-tuple costs include the cost of evaluating the lefthand
|
|
* expressions, plus the cost of probing the hashtable. We already
|
|
* accounted for the lefthand expressions as part of the testexpr, and
|
|
* will also have counted one cpu_operator_cost for each comparison
|
|
* operator. That is probably too low for the probing cost, but it's
|
|
* hard to make a better estimate, so live with it for now.
|
|
*/
|
|
}
|
|
else
|
|
{
|
|
/*
|
|
* Otherwise we will be rescanning the subplan output on each
|
|
* evaluation. We need to estimate how much of the output we will
|
|
* actually need to scan. NOTE: this logic should agree with the
|
|
* tuple_fraction estimates used by make_subplan() in
|
|
* plan/subselect.c.
|
|
*/
|
|
Cost plan_run_cost = plan->total_cost - plan->startup_cost;
|
|
|
|
if (subplan->subLinkType == EXISTS_SUBLINK)
|
|
{
|
|
/* we only need to fetch 1 tuple; clamp to avoid zero divide */
|
|
sp_cost.per_tuple += plan_run_cost / clamp_row_est(plan->plan_rows);
|
|
}
|
|
else if (subplan->subLinkType == ALL_SUBLINK ||
|
|
subplan->subLinkType == ANY_SUBLINK)
|
|
{
|
|
/* assume we need 50% of the tuples */
|
|
sp_cost.per_tuple += 0.50 * plan_run_cost;
|
|
/* also charge a cpu_operator_cost per row examined */
|
|
sp_cost.per_tuple += 0.50 * plan->plan_rows * cpu_operator_cost;
|
|
}
|
|
else
|
|
{
|
|
/* assume we need all tuples */
|
|
sp_cost.per_tuple += plan_run_cost;
|
|
}
|
|
|
|
/*
|
|
* Also account for subplan's startup cost. If the subplan is
|
|
* uncorrelated or undirect correlated, AND its topmost node is one
|
|
* that materializes its output, assume that we'll only need to pay
|
|
* its startup cost once; otherwise assume we pay the startup cost
|
|
* every time.
|
|
*/
|
|
if (subplan->parParam == NIL &&
|
|
ExecMaterializesOutput(nodeTag(plan)))
|
|
sp_cost.startup += plan->startup_cost;
|
|
else
|
|
sp_cost.per_tuple += plan->startup_cost;
|
|
}
|
|
|
|
subplan->startup_cost = sp_cost.startup;
|
|
subplan->per_call_cost = sp_cost.per_tuple;
|
|
}
|
|
|
|
|
|
/*
|
|
* cost_rescan
|
|
* Given a finished Path, estimate the costs of rescanning it after
|
|
* having done so the first time. For some Path types a rescan is
|
|
* cheaper than an original scan (if no parameters change), and this
|
|
* function embodies knowledge about that. The default is to return
|
|
* the same costs stored in the Path. (Note that the cost estimates
|
|
* actually stored in Paths are always for first scans.)
|
|
*
|
|
* This function is not currently intended to model effects such as rescans
|
|
* being cheaper due to disk block caching; what we are concerned with is
|
|
* plan types wherein the executor caches results explicitly, or doesn't
|
|
* redo startup calculations, etc.
|
|
*/
|
|
static void
|
|
cost_rescan(PlannerInfo *root, Path *path,
|
|
Cost *rescan_startup_cost, /* output parameters */
|
|
Cost *rescan_total_cost)
|
|
{
|
|
switch (path->pathtype)
|
|
{
|
|
case T_FunctionScan:
|
|
|
|
/*
|
|
* Currently, nodeFunctionscan.c always executes the function to
|
|
* completion before returning any rows, and caches the results in
|
|
* a tuplestore. So the function eval cost is all startup cost
|
|
* and isn't paid over again on rescans. However, all run costs
|
|
* will be paid over again.
|
|
*/
|
|
*rescan_startup_cost = 0;
|
|
*rescan_total_cost = path->total_cost - path->startup_cost;
|
|
break;
|
|
case T_HashJoin:
|
|
|
|
/*
|
|
* If it's a single-batch join, we don't need to rebuild the hash
|
|
* table during a rescan.
|
|
*/
|
|
if (((HashPath *) path)->num_batches == 1)
|
|
{
|
|
/* Startup cost is exactly the cost of hash table building */
|
|
*rescan_startup_cost = 0;
|
|
*rescan_total_cost = path->total_cost - path->startup_cost;
|
|
}
|
|
else
|
|
{
|
|
/* Otherwise, no special treatment */
|
|
*rescan_startup_cost = path->startup_cost;
|
|
*rescan_total_cost = path->total_cost;
|
|
}
|
|
break;
|
|
case T_CteScan:
|
|
case T_WorkTableScan:
|
|
{
|
|
/*
|
|
* These plan types materialize their final result in a
|
|
* tuplestore or tuplesort object. So the rescan cost is only
|
|
* cpu_tuple_cost per tuple, unless the result is large enough
|
|
* to spill to disk.
|
|
*/
|
|
Cost run_cost = cpu_tuple_cost * path->rows;
|
|
double nbytes = relation_byte_size(path->rows,
|
|
path->pathtarget->width);
|
|
long work_mem_bytes = work_mem * 1024L;
|
|
|
|
if (nbytes > work_mem_bytes)
|
|
{
|
|
/* It will spill, so account for re-read cost */
|
|
double npages = ceil(nbytes / BLCKSZ);
|
|
|
|
run_cost += seq_page_cost * npages;
|
|
}
|
|
*rescan_startup_cost = 0;
|
|
*rescan_total_cost = run_cost;
|
|
}
|
|
break;
|
|
case T_Material:
|
|
case T_Sort:
|
|
{
|
|
/*
|
|
* These plan types not only materialize their results, but do
|
|
* not implement qual filtering or projection. So they are
|
|
* even cheaper to rescan than the ones above. We charge only
|
|
* cpu_operator_cost per tuple. (Note: keep that in sync with
|
|
* the run_cost charge in cost_sort, and also see comments in
|
|
* cost_material before you change it.)
|
|
*/
|
|
Cost run_cost = cpu_operator_cost * path->rows;
|
|
double nbytes = relation_byte_size(path->rows,
|
|
path->pathtarget->width);
|
|
long work_mem_bytes = work_mem * 1024L;
|
|
|
|
if (nbytes > work_mem_bytes)
|
|
{
|
|
/* It will spill, so account for re-read cost */
|
|
double npages = ceil(nbytes / BLCKSZ);
|
|
|
|
run_cost += seq_page_cost * npages;
|
|
}
|
|
*rescan_startup_cost = 0;
|
|
*rescan_total_cost = run_cost;
|
|
}
|
|
break;
|
|
case T_Memoize:
|
|
/* All the hard work is done by cost_memoize_rescan */
|
|
cost_memoize_rescan(root, (MemoizePath *) path,
|
|
rescan_startup_cost, rescan_total_cost);
|
|
break;
|
|
default:
|
|
*rescan_startup_cost = path->startup_cost;
|
|
*rescan_total_cost = path->total_cost;
|
|
break;
|
|
}
|
|
}
|
|
|
|
|
|
/*
|
|
* cost_qual_eval
|
|
* Estimate the CPU costs of evaluating a WHERE clause.
|
|
* The input can be either an implicitly-ANDed list of boolean
|
|
* expressions, or a list of RestrictInfo nodes. (The latter is
|
|
* preferred since it allows caching of the results.)
|
|
* The result includes both a one-time (startup) component,
|
|
* and a per-evaluation component.
|
|
*/
|
|
void
|
|
cost_qual_eval(QualCost *cost, List *quals, PlannerInfo *root)
|
|
{
|
|
cost_qual_eval_context context;
|
|
ListCell *l;
|
|
|
|
context.root = root;
|
|
context.total.startup = 0;
|
|
context.total.per_tuple = 0;
|
|
|
|
/* We don't charge any cost for the implicit ANDing at top level ... */
|
|
|
|
foreach(l, quals)
|
|
{
|
|
Node *qual = (Node *) lfirst(l);
|
|
|
|
cost_qual_eval_walker(qual, &context);
|
|
}
|
|
|
|
*cost = context.total;
|
|
}
|
|
|
|
/*
|
|
* cost_qual_eval_node
|
|
* As above, for a single RestrictInfo or expression.
|
|
*/
|
|
void
|
|
cost_qual_eval_node(QualCost *cost, Node *qual, PlannerInfo *root)
|
|
{
|
|
cost_qual_eval_context context;
|
|
|
|
context.root = root;
|
|
context.total.startup = 0;
|
|
context.total.per_tuple = 0;
|
|
|
|
cost_qual_eval_walker(qual, &context);
|
|
|
|
*cost = context.total;
|
|
}
|
|
|
|
static bool
|
|
cost_qual_eval_walker(Node *node, cost_qual_eval_context *context)
|
|
{
|
|
if (node == NULL)
|
|
return false;
|
|
|
|
/*
|
|
* RestrictInfo nodes contain an eval_cost field reserved for this
|
|
* routine's use, so that it's not necessary to evaluate the qual clause's
|
|
* cost more than once. If the clause's cost hasn't been computed yet,
|
|
* the field's startup value will contain -1.
|
|
*/
|
|
if (IsA(node, RestrictInfo))
|
|
{
|
|
RestrictInfo *rinfo = (RestrictInfo *) node;
|
|
|
|
if (rinfo->eval_cost.startup < 0)
|
|
{
|
|
cost_qual_eval_context locContext;
|
|
|
|
locContext.root = context->root;
|
|
locContext.total.startup = 0;
|
|
locContext.total.per_tuple = 0;
|
|
|
|
/*
|
|
* For an OR clause, recurse into the marked-up tree so that we
|
|
* set the eval_cost for contained RestrictInfos too.
|
|
*/
|
|
if (rinfo->orclause)
|
|
cost_qual_eval_walker((Node *) rinfo->orclause, &locContext);
|
|
else
|
|
cost_qual_eval_walker((Node *) rinfo->clause, &locContext);
|
|
|
|
/*
|
|
* If the RestrictInfo is marked pseudoconstant, it will be tested
|
|
* only once, so treat its cost as all startup cost.
|
|
*/
|
|
if (rinfo->pseudoconstant)
|
|
{
|
|
/* count one execution during startup */
|
|
locContext.total.startup += locContext.total.per_tuple;
|
|
locContext.total.per_tuple = 0;
|
|
}
|
|
rinfo->eval_cost = locContext.total;
|
|
}
|
|
context->total.startup += rinfo->eval_cost.startup;
|
|
context->total.per_tuple += rinfo->eval_cost.per_tuple;
|
|
/* do NOT recurse into children */
|
|
return false;
|
|
}
|
|
|
|
/*
|
|
* For each operator or function node in the given tree, we charge the
|
|
* estimated execution cost given by pg_proc.procost (remember to multiply
|
|
* this by cpu_operator_cost).
|
|
*
|
|
* Vars and Consts are charged zero, and so are boolean operators (AND,
|
|
* OR, NOT). Simplistic, but a lot better than no model at all.
|
|
*
|
|
* Should we try to account for the possibility of short-circuit
|
|
* evaluation of AND/OR? Probably *not*, because that would make the
|
|
* results depend on the clause ordering, and we are not in any position
|
|
* to expect that the current ordering of the clauses is the one that's
|
|
* going to end up being used. The above per-RestrictInfo caching would
|
|
* not mix well with trying to re-order clauses anyway.
|
|
*
|
|
* Another issue that is entirely ignored here is that if a set-returning
|
|
* function is below top level in the tree, the functions/operators above
|
|
* it will need to be evaluated multiple times. In practical use, such
|
|
* cases arise so seldom as to not be worth the added complexity needed;
|
|
* moreover, since our rowcount estimates for functions tend to be pretty
|
|
* phony, the results would also be pretty phony.
|
|
*/
|
|
if (IsA(node, FuncExpr))
|
|
{
|
|
add_function_cost(context->root, ((FuncExpr *) node)->funcid, node,
|
|
&context->total);
|
|
}
|
|
else if (IsA(node, OpExpr) ||
|
|
IsA(node, DistinctExpr) ||
|
|
IsA(node, NullIfExpr))
|
|
{
|
|
/* rely on struct equivalence to treat these all alike */
|
|
set_opfuncid((OpExpr *) node);
|
|
add_function_cost(context->root, ((OpExpr *) node)->opfuncid, node,
|
|
&context->total);
|
|
}
|
|
else if (IsA(node, ScalarArrayOpExpr))
|
|
{
|
|
ScalarArrayOpExpr *saop = (ScalarArrayOpExpr *) node;
|
|
Node *arraynode = (Node *) lsecond(saop->args);
|
|
QualCost sacosts;
|
|
QualCost hcosts;
|
|
int estarraylen = estimate_array_length(arraynode);
|
|
|
|
set_sa_opfuncid(saop);
|
|
sacosts.startup = sacosts.per_tuple = 0;
|
|
add_function_cost(context->root, saop->opfuncid, NULL,
|
|
&sacosts);
|
|
|
|
if (OidIsValid(saop->hashfuncid))
|
|
{
|
|
/* Handle costs for hashed ScalarArrayOpExpr */
|
|
hcosts.startup = hcosts.per_tuple = 0;
|
|
|
|
add_function_cost(context->root, saop->hashfuncid, NULL, &hcosts);
|
|
context->total.startup += sacosts.startup + hcosts.startup;
|
|
|
|
/* Estimate the cost of building the hashtable. */
|
|
context->total.startup += estarraylen * hcosts.per_tuple;
|
|
|
|
/*
|
|
* XXX should we charge a little bit for sacosts.per_tuple when
|
|
* building the table, or is it ok to assume there will be zero
|
|
* hash collision?
|
|
*/
|
|
|
|
/*
|
|
* Charge for hashtable lookups. Charge a single hash and a
|
|
* single comparison.
|
|
*/
|
|
context->total.per_tuple += hcosts.per_tuple + sacosts.per_tuple;
|
|
}
|
|
else
|
|
{
|
|
/*
|
|
* Estimate that the operator will be applied to about half of the
|
|
* array elements before the answer is determined.
|
|
*/
|
|
context->total.startup += sacosts.startup;
|
|
context->total.per_tuple += sacosts.per_tuple *
|
|
estimate_array_length(arraynode) * 0.5;
|
|
}
|
|
}
|
|
else if (IsA(node, Aggref) ||
|
|
IsA(node, WindowFunc))
|
|
{
|
|
/*
|
|
* Aggref and WindowFunc nodes are (and should be) treated like Vars,
|
|
* ie, zero execution cost in the current model, because they behave
|
|
* essentially like Vars at execution. We disregard the costs of
|
|
* their input expressions for the same reason. The actual execution
|
|
* costs of the aggregate/window functions and their arguments have to
|
|
* be factored into plan-node-specific costing of the Agg or WindowAgg
|
|
* plan node.
|
|
*/
|
|
return false; /* don't recurse into children */
|
|
}
|
|
else if (IsA(node, GroupingFunc))
|
|
{
|
|
/* Treat this as having cost 1 */
|
|
context->total.per_tuple += cpu_operator_cost;
|
|
return false; /* don't recurse into children */
|
|
}
|
|
else if (IsA(node, CoerceViaIO))
|
|
{
|
|
CoerceViaIO *iocoerce = (CoerceViaIO *) node;
|
|
Oid iofunc;
|
|
Oid typioparam;
|
|
bool typisvarlena;
|
|
|
|
/* check the result type's input function */
|
|
getTypeInputInfo(iocoerce->resulttype,
|
|
&iofunc, &typioparam);
|
|
add_function_cost(context->root, iofunc, NULL,
|
|
&context->total);
|
|
/* check the input type's output function */
|
|
getTypeOutputInfo(exprType((Node *) iocoerce->arg),
|
|
&iofunc, &typisvarlena);
|
|
add_function_cost(context->root, iofunc, NULL,
|
|
&context->total);
|
|
}
|
|
else if (IsA(node, ArrayCoerceExpr))
|
|
{
|
|
ArrayCoerceExpr *acoerce = (ArrayCoerceExpr *) node;
|
|
QualCost perelemcost;
|
|
|
|
cost_qual_eval_node(&perelemcost, (Node *) acoerce->elemexpr,
|
|
context->root);
|
|
context->total.startup += perelemcost.startup;
|
|
if (perelemcost.per_tuple > 0)
|
|
context->total.per_tuple += perelemcost.per_tuple *
|
|
estimate_array_length((Node *) acoerce->arg);
|
|
}
|
|
else if (IsA(node, RowCompareExpr))
|
|
{
|
|
/* Conservatively assume we will check all the columns */
|
|
RowCompareExpr *rcexpr = (RowCompareExpr *) node;
|
|
ListCell *lc;
|
|
|
|
foreach(lc, rcexpr->opnos)
|
|
{
|
|
Oid opid = lfirst_oid(lc);
|
|
|
|
add_function_cost(context->root, get_opcode(opid), NULL,
|
|
&context->total);
|
|
}
|
|
}
|
|
else if (IsA(node, MinMaxExpr) ||
|
|
IsA(node, SQLValueFunction) ||
|
|
IsA(node, XmlExpr) ||
|
|
IsA(node, CoerceToDomain) ||
|
|
IsA(node, NextValueExpr))
|
|
{
|
|
/* Treat all these as having cost 1 */
|
|
context->total.per_tuple += cpu_operator_cost;
|
|
}
|
|
else if (IsA(node, CurrentOfExpr))
|
|
{
|
|
/* Report high cost to prevent selection of anything but TID scan */
|
|
context->total.startup += disable_cost;
|
|
}
|
|
else if (IsA(node, SubLink))
|
|
{
|
|
/* This routine should not be applied to un-planned expressions */
|
|
elog(ERROR, "cannot handle unplanned sub-select");
|
|
}
|
|
else if (IsA(node, SubPlan))
|
|
{
|
|
/*
|
|
* A subplan node in an expression typically indicates that the
|
|
* subplan will be executed on each evaluation, so charge accordingly.
|
|
* (Sub-selects that can be executed as InitPlans have already been
|
|
* removed from the expression.)
|
|
*/
|
|
SubPlan *subplan = (SubPlan *) node;
|
|
|
|
context->total.startup += subplan->startup_cost;
|
|
context->total.per_tuple += subplan->per_call_cost;
|
|
|
|
/*
|
|
* We don't want to recurse into the testexpr, because it was already
|
|
* counted in the SubPlan node's costs. So we're done.
|
|
*/
|
|
return false;
|
|
}
|
|
else if (IsA(node, AlternativeSubPlan))
|
|
{
|
|
/*
|
|
* Arbitrarily use the first alternative plan for costing. (We should
|
|
* certainly only include one alternative, and we don't yet have
|
|
* enough information to know which one the executor is most likely to
|
|
* use.)
|
|
*/
|
|
AlternativeSubPlan *asplan = (AlternativeSubPlan *) node;
|
|
|
|
return cost_qual_eval_walker((Node *) linitial(asplan->subplans),
|
|
context);
|
|
}
|
|
else if (IsA(node, PlaceHolderVar))
|
|
{
|
|
/*
|
|
* A PlaceHolderVar should be given cost zero when considering general
|
|
* expression evaluation costs. The expense of doing the contained
|
|
* expression is charged as part of the tlist eval costs of the scan
|
|
* or join where the PHV is first computed (see set_rel_width and
|
|
* add_placeholders_to_joinrel). If we charged it again here, we'd be
|
|
* double-counting the cost for each level of plan that the PHV
|
|
* bubbles up through. Hence, return without recursing into the
|
|
* phexpr.
|
|
*/
|
|
return false;
|
|
}
|
|
|
|
/* recurse into children */
|
|
return expression_tree_walker(node, cost_qual_eval_walker,
|
|
(void *) context);
|
|
}
|
|
|
|
/*
|
|
* get_restriction_qual_cost
|
|
* Compute evaluation costs of a baserel's restriction quals, plus any
|
|
* movable join quals that have been pushed down to the scan.
|
|
* Results are returned into *qpqual_cost.
|
|
*
|
|
* This is a convenience subroutine that works for seqscans and other cases
|
|
* where all the given quals will be evaluated the hard way. It's not useful
|
|
* for cost_index(), for example, where the index machinery takes care of
|
|
* some of the quals. We assume baserestrictcost was previously set by
|
|
* set_baserel_size_estimates().
|
|
*/
|
|
static void
|
|
get_restriction_qual_cost(PlannerInfo *root, RelOptInfo *baserel,
|
|
ParamPathInfo *param_info,
|
|
QualCost *qpqual_cost)
|
|
{
|
|
if (param_info)
|
|
{
|
|
/* Include costs of pushed-down clauses */
|
|
cost_qual_eval(qpqual_cost, param_info->ppi_clauses, root);
|
|
|
|
qpqual_cost->startup += baserel->baserestrictcost.startup;
|
|
qpqual_cost->per_tuple += baserel->baserestrictcost.per_tuple;
|
|
}
|
|
else
|
|
*qpqual_cost = baserel->baserestrictcost;
|
|
}
|
|
|
|
|
|
/*
|
|
* compute_semi_anti_join_factors
|
|
* Estimate how much of the inner input a SEMI, ANTI, or inner_unique join
|
|
* can be expected to scan.
|
|
*
|
|
* In a hash or nestloop SEMI/ANTI join, the executor will stop scanning
|
|
* inner rows as soon as it finds a match to the current outer row.
|
|
* The same happens if we have detected the inner rel is unique.
|
|
* We should therefore adjust some of the cost components for this effect.
|
|
* This function computes some estimates needed for these adjustments.
|
|
* These estimates will be the same regardless of the particular paths used
|
|
* for the outer and inner relation, so we compute these once and then pass
|
|
* them to all the join cost estimation functions.
|
|
*
|
|
* Input parameters:
|
|
* joinrel: join relation under consideration
|
|
* outerrel: outer relation under consideration
|
|
* innerrel: inner relation under consideration
|
|
* jointype: if not JOIN_SEMI or JOIN_ANTI, we assume it's inner_unique
|
|
* sjinfo: SpecialJoinInfo relevant to this join
|
|
* restrictlist: join quals
|
|
* Output parameters:
|
|
* *semifactors is filled in (see pathnodes.h for field definitions)
|
|
*/
|
|
void
|
|
compute_semi_anti_join_factors(PlannerInfo *root,
|
|
RelOptInfo *joinrel,
|
|
RelOptInfo *outerrel,
|
|
RelOptInfo *innerrel,
|
|
JoinType jointype,
|
|
SpecialJoinInfo *sjinfo,
|
|
List *restrictlist,
|
|
SemiAntiJoinFactors *semifactors)
|
|
{
|
|
Selectivity jselec;
|
|
Selectivity nselec;
|
|
Selectivity avgmatch;
|
|
SpecialJoinInfo norm_sjinfo;
|
|
List *joinquals;
|
|
ListCell *l;
|
|
|
|
/*
|
|
* In an ANTI join, we must ignore clauses that are "pushed down", since
|
|
* those won't affect the match logic. In a SEMI join, we do not
|
|
* distinguish joinquals from "pushed down" quals, so just use the whole
|
|
* restrictinfo list. For other outer join types, we should consider only
|
|
* non-pushed-down quals, so that this devolves to an IS_OUTER_JOIN check.
|
|
*/
|
|
if (IS_OUTER_JOIN(jointype))
|
|
{
|
|
joinquals = NIL;
|
|
foreach(l, restrictlist)
|
|
{
|
|
RestrictInfo *rinfo = lfirst_node(RestrictInfo, l);
|
|
|
|
if (!RINFO_IS_PUSHED_DOWN(rinfo, joinrel->relids))
|
|
joinquals = lappend(joinquals, rinfo);
|
|
}
|
|
}
|
|
else
|
|
joinquals = restrictlist;
|
|
|
|
/*
|
|
* Get the JOIN_SEMI or JOIN_ANTI selectivity of the join clauses.
|
|
*/
|
|
jselec = clauselist_selectivity(root,
|
|
joinquals,
|
|
0,
|
|
(jointype == JOIN_ANTI) ? JOIN_ANTI : JOIN_SEMI,
|
|
sjinfo);
|
|
|
|
/*
|
|
* Also get the normal inner-join selectivity of the join clauses.
|
|
*/
|
|
norm_sjinfo.type = T_SpecialJoinInfo;
|
|
norm_sjinfo.min_lefthand = outerrel->relids;
|
|
norm_sjinfo.min_righthand = innerrel->relids;
|
|
norm_sjinfo.syn_lefthand = outerrel->relids;
|
|
norm_sjinfo.syn_righthand = innerrel->relids;
|
|
norm_sjinfo.jointype = JOIN_INNER;
|
|
norm_sjinfo.ojrelid = 0;
|
|
norm_sjinfo.commute_above_l = NULL;
|
|
norm_sjinfo.commute_above_r = NULL;
|
|
norm_sjinfo.commute_below_l = NULL;
|
|
norm_sjinfo.commute_below_r = NULL;
|
|
/* we don't bother trying to make the remaining fields valid */
|
|
norm_sjinfo.lhs_strict = false;
|
|
norm_sjinfo.semi_can_btree = false;
|
|
norm_sjinfo.semi_can_hash = false;
|
|
norm_sjinfo.semi_operators = NIL;
|
|
norm_sjinfo.semi_rhs_exprs = NIL;
|
|
|
|
nselec = clauselist_selectivity(root,
|
|
joinquals,
|
|
0,
|
|
JOIN_INNER,
|
|
&norm_sjinfo);
|
|
|
|
/* Avoid leaking a lot of ListCells */
|
|
if (IS_OUTER_JOIN(jointype))
|
|
list_free(joinquals);
|
|
|
|
/*
|
|
* jselec can be interpreted as the fraction of outer-rel rows that have
|
|
* any matches (this is true for both SEMI and ANTI cases). And nselec is
|
|
* the fraction of the Cartesian product that matches. So, the average
|
|
* number of matches for each outer-rel row that has at least one match is
|
|
* nselec * inner_rows / jselec.
|
|
*
|
|
* Note: it is correct to use the inner rel's "rows" count here, even
|
|
* though we might later be considering a parameterized inner path with
|
|
* fewer rows. This is because we have included all the join clauses in
|
|
* the selectivity estimate.
|
|
*/
|
|
if (jselec > 0) /* protect against zero divide */
|
|
{
|
|
avgmatch = nselec * innerrel->rows / jselec;
|
|
/* Clamp to sane range */
|
|
avgmatch = Max(1.0, avgmatch);
|
|
}
|
|
else
|
|
avgmatch = 1.0;
|
|
|
|
semifactors->outer_match_frac = jselec;
|
|
semifactors->match_count = avgmatch;
|
|
}
|
|
|
|
/*
|
|
* has_indexed_join_quals
|
|
* Check whether all the joinquals of a nestloop join are used as
|
|
* inner index quals.
|
|
*
|
|
* If the inner path of a SEMI/ANTI join is an indexscan (including bitmap
|
|
* indexscan) that uses all the joinquals as indexquals, we can assume that an
|
|
* unmatched outer tuple is cheap to process, whereas otherwise it's probably
|
|
* expensive.
|
|
*/
|
|
static bool
|
|
has_indexed_join_quals(NestPath *path)
|
|
{
|
|
JoinPath *joinpath = &path->jpath;
|
|
Relids joinrelids = joinpath->path.parent->relids;
|
|
Path *innerpath = joinpath->innerjoinpath;
|
|
List *indexclauses;
|
|
bool found_one;
|
|
ListCell *lc;
|
|
|
|
/* If join still has quals to evaluate, it's not fast */
|
|
if (joinpath->joinrestrictinfo != NIL)
|
|
return false;
|
|
/* Nor if the inner path isn't parameterized at all */
|
|
if (innerpath->param_info == NULL)
|
|
return false;
|
|
|
|
/* Find the indexclauses list for the inner scan */
|
|
switch (innerpath->pathtype)
|
|
{
|
|
case T_IndexScan:
|
|
case T_IndexOnlyScan:
|
|
indexclauses = ((IndexPath *) innerpath)->indexclauses;
|
|
break;
|
|
case T_BitmapHeapScan:
|
|
{
|
|
/* Accept only a simple bitmap scan, not AND/OR cases */
|
|
Path *bmqual = ((BitmapHeapPath *) innerpath)->bitmapqual;
|
|
|
|
if (IsA(bmqual, IndexPath))
|
|
indexclauses = ((IndexPath *) bmqual)->indexclauses;
|
|
else
|
|
return false;
|
|
break;
|
|
}
|
|
default:
|
|
|
|
/*
|
|
* If it's not a simple indexscan, it probably doesn't run quickly
|
|
* for zero rows out, even if it's a parameterized path using all
|
|
* the joinquals.
|
|
*/
|
|
return false;
|
|
}
|
|
|
|
/*
|
|
* Examine the inner path's param clauses. Any that are from the outer
|
|
* path must be found in the indexclauses list, either exactly or in an
|
|
* equivalent form generated by equivclass.c. Also, we must find at least
|
|
* one such clause, else it's a clauseless join which isn't fast.
|
|
*/
|
|
found_one = false;
|
|
foreach(lc, innerpath->param_info->ppi_clauses)
|
|
{
|
|
RestrictInfo *rinfo = (RestrictInfo *) lfirst(lc);
|
|
|
|
if (join_clause_is_movable_into(rinfo,
|
|
innerpath->parent->relids,
|
|
joinrelids))
|
|
{
|
|
if (!is_redundant_with_indexclauses(rinfo, indexclauses))
|
|
return false;
|
|
found_one = true;
|
|
}
|
|
}
|
|
return found_one;
|
|
}
|
|
|
|
|
|
/*
|
|
* approx_tuple_count
|
|
* Quick-and-dirty estimation of the number of join rows passing
|
|
* a set of qual conditions.
|
|
*
|
|
* The quals can be either an implicitly-ANDed list of boolean expressions,
|
|
* or a list of RestrictInfo nodes (typically the latter).
|
|
*
|
|
* We intentionally compute the selectivity under JOIN_INNER rules, even
|
|
* if it's some type of outer join. This is appropriate because we are
|
|
* trying to figure out how many tuples pass the initial merge or hash
|
|
* join step.
|
|
*
|
|
* This is quick-and-dirty because we bypass clauselist_selectivity, and
|
|
* simply multiply the independent clause selectivities together. Now
|
|
* clauselist_selectivity often can't do any better than that anyhow, but
|
|
* for some situations (such as range constraints) it is smarter. However,
|
|
* we can't effectively cache the results of clauselist_selectivity, whereas
|
|
* the individual clause selectivities can be and are cached.
|
|
*
|
|
* Since we are only using the results to estimate how many potential
|
|
* output tuples are generated and passed through qpqual checking, it
|
|
* seems OK to live with the approximation.
|
|
*/
|
|
static double
|
|
approx_tuple_count(PlannerInfo *root, JoinPath *path, List *quals)
|
|
{
|
|
double tuples;
|
|
double outer_tuples = path->outerjoinpath->rows;
|
|
double inner_tuples = path->innerjoinpath->rows;
|
|
SpecialJoinInfo sjinfo;
|
|
Selectivity selec = 1.0;
|
|
ListCell *l;
|
|
|
|
/*
|
|
* Make up a SpecialJoinInfo for JOIN_INNER semantics.
|
|
*/
|
|
sjinfo.type = T_SpecialJoinInfo;
|
|
sjinfo.min_lefthand = path->outerjoinpath->parent->relids;
|
|
sjinfo.min_righthand = path->innerjoinpath->parent->relids;
|
|
sjinfo.syn_lefthand = path->outerjoinpath->parent->relids;
|
|
sjinfo.syn_righthand = path->innerjoinpath->parent->relids;
|
|
sjinfo.jointype = JOIN_INNER;
|
|
sjinfo.ojrelid = 0;
|
|
sjinfo.commute_above_l = NULL;
|
|
sjinfo.commute_above_r = NULL;
|
|
sjinfo.commute_below_l = NULL;
|
|
sjinfo.commute_below_r = NULL;
|
|
/* we don't bother trying to make the remaining fields valid */
|
|
sjinfo.lhs_strict = false;
|
|
sjinfo.semi_can_btree = false;
|
|
sjinfo.semi_can_hash = false;
|
|
sjinfo.semi_operators = NIL;
|
|
sjinfo.semi_rhs_exprs = NIL;
|
|
|
|
/* Get the approximate selectivity */
|
|
foreach(l, quals)
|
|
{
|
|
Node *qual = (Node *) lfirst(l);
|
|
|
|
/* Note that clause_selectivity will be able to cache its result */
|
|
selec *= clause_selectivity(root, qual, 0, JOIN_INNER, &sjinfo);
|
|
}
|
|
|
|
/* Apply it to the input relation sizes */
|
|
tuples = selec * outer_tuples * inner_tuples;
|
|
|
|
return clamp_row_est(tuples);
|
|
}
|
|
|
|
|
|
/*
|
|
* set_baserel_size_estimates
|
|
* Set the size estimates for the given base relation.
|
|
*
|
|
* The rel's targetlist and restrictinfo list must have been constructed
|
|
* already, and rel->tuples must be set.
|
|
*
|
|
* We set the following fields of the rel node:
|
|
* rows: the estimated number of output tuples (after applying
|
|
* restriction clauses).
|
|
* width: the estimated average output tuple width in bytes.
|
|
* baserestrictcost: estimated cost of evaluating baserestrictinfo clauses.
|
|
*/
|
|
void
|
|
set_baserel_size_estimates(PlannerInfo *root, RelOptInfo *rel)
|
|
{
|
|
double nrows;
|
|
|
|
/* Should only be applied to base relations */
|
|
Assert(rel->relid > 0);
|
|
|
|
nrows = rel->tuples *
|
|
clauselist_selectivity(root,
|
|
rel->baserestrictinfo,
|
|
0,
|
|
JOIN_INNER,
|
|
NULL);
|
|
|
|
rel->rows = clamp_row_est(nrows);
|
|
|
|
cost_qual_eval(&rel->baserestrictcost, rel->baserestrictinfo, root);
|
|
|
|
set_rel_width(root, rel);
|
|
}
|
|
|
|
/*
|
|
* get_parameterized_baserel_size
|
|
* Make a size estimate for a parameterized scan of a base relation.
|
|
*
|
|
* 'param_clauses' lists the additional join clauses to be used.
|
|
*
|
|
* set_baserel_size_estimates must have been applied already.
|
|
*/
|
|
double
|
|
get_parameterized_baserel_size(PlannerInfo *root, RelOptInfo *rel,
|
|
List *param_clauses)
|
|
{
|
|
List *allclauses;
|
|
double nrows;
|
|
|
|
/*
|
|
* Estimate the number of rows returned by the parameterized scan, knowing
|
|
* that it will apply all the extra join clauses as well as the rel's own
|
|
* restriction clauses. Note that we force the clauses to be treated as
|
|
* non-join clauses during selectivity estimation.
|
|
*/
|
|
allclauses = list_concat_copy(param_clauses, rel->baserestrictinfo);
|
|
nrows = rel->tuples *
|
|
clauselist_selectivity(root,
|
|
allclauses,
|
|
rel->relid, /* do not use 0! */
|
|
JOIN_INNER,
|
|
NULL);
|
|
nrows = clamp_row_est(nrows);
|
|
/* For safety, make sure result is not more than the base estimate */
|
|
if (nrows > rel->rows)
|
|
nrows = rel->rows;
|
|
return nrows;
|
|
}
|
|
|
|
/*
|
|
* set_joinrel_size_estimates
|
|
* Set the size estimates for the given join relation.
|
|
*
|
|
* The rel's targetlist must have been constructed already, and a
|
|
* restriction clause list that matches the given component rels must
|
|
* be provided.
|
|
*
|
|
* Since there is more than one way to make a joinrel for more than two
|
|
* base relations, the results we get here could depend on which component
|
|
* rel pair is provided. In theory we should get the same answers no matter
|
|
* which pair is provided; in practice, since the selectivity estimation
|
|
* routines don't handle all cases equally well, we might not. But there's
|
|
* not much to be done about it. (Would it make sense to repeat the
|
|
* calculations for each pair of input rels that's encountered, and somehow
|
|
* average the results? Probably way more trouble than it's worth, and
|
|
* anyway we must keep the rowcount estimate the same for all paths for the
|
|
* joinrel.)
|
|
*
|
|
* We set only the rows field here. The reltarget field was already set by
|
|
* build_joinrel_tlist, and baserestrictcost is not used for join rels.
|
|
*/
|
|
void
|
|
set_joinrel_size_estimates(PlannerInfo *root, RelOptInfo *rel,
|
|
RelOptInfo *outer_rel,
|
|
RelOptInfo *inner_rel,
|
|
SpecialJoinInfo *sjinfo,
|
|
List *restrictlist)
|
|
{
|
|
rel->rows = calc_joinrel_size_estimate(root,
|
|
rel,
|
|
outer_rel,
|
|
inner_rel,
|
|
outer_rel->rows,
|
|
inner_rel->rows,
|
|
sjinfo,
|
|
restrictlist);
|
|
}
|
|
|
|
/*
|
|
* get_parameterized_joinrel_size
|
|
* Make a size estimate for a parameterized scan of a join relation.
|
|
*
|
|
* 'rel' is the joinrel under consideration.
|
|
* 'outer_path', 'inner_path' are (probably also parameterized) Paths that
|
|
* produce the relations being joined.
|
|
* 'sjinfo' is any SpecialJoinInfo relevant to this join.
|
|
* 'restrict_clauses' lists the join clauses that need to be applied at the
|
|
* join node (including any movable clauses that were moved down to this join,
|
|
* and not including any movable clauses that were pushed down into the
|
|
* child paths).
|
|
*
|
|
* set_joinrel_size_estimates must have been applied already.
|
|
*/
|
|
double
|
|
get_parameterized_joinrel_size(PlannerInfo *root, RelOptInfo *rel,
|
|
Path *outer_path,
|
|
Path *inner_path,
|
|
SpecialJoinInfo *sjinfo,
|
|
List *restrict_clauses)
|
|
{
|
|
double nrows;
|
|
|
|
/*
|
|
* Estimate the number of rows returned by the parameterized join as the
|
|
* sizes of the input paths times the selectivity of the clauses that have
|
|
* ended up at this join node.
|
|
*
|
|
* As with set_joinrel_size_estimates, the rowcount estimate could depend
|
|
* on the pair of input paths provided, though ideally we'd get the same
|
|
* estimate for any pair with the same parameterization.
|
|
*/
|
|
nrows = calc_joinrel_size_estimate(root,
|
|
rel,
|
|
outer_path->parent,
|
|
inner_path->parent,
|
|
outer_path->rows,
|
|
inner_path->rows,
|
|
sjinfo,
|
|
restrict_clauses);
|
|
/* For safety, make sure result is not more than the base estimate */
|
|
if (nrows > rel->rows)
|
|
nrows = rel->rows;
|
|
return nrows;
|
|
}
|
|
|
|
/*
|
|
* calc_joinrel_size_estimate
|
|
* Workhorse for set_joinrel_size_estimates and
|
|
* get_parameterized_joinrel_size.
|
|
*
|
|
* outer_rel/inner_rel are the relations being joined, but they should be
|
|
* assumed to have sizes outer_rows/inner_rows; those numbers might be less
|
|
* than what rel->rows says, when we are considering parameterized paths.
|
|
*/
|
|
static double
|
|
calc_joinrel_size_estimate(PlannerInfo *root,
|
|
RelOptInfo *joinrel,
|
|
RelOptInfo *outer_rel,
|
|
RelOptInfo *inner_rel,
|
|
double outer_rows,
|
|
double inner_rows,
|
|
SpecialJoinInfo *sjinfo,
|
|
List *restrictlist)
|
|
{
|
|
JoinType jointype = sjinfo->jointype;
|
|
Selectivity fkselec;
|
|
Selectivity jselec;
|
|
Selectivity pselec;
|
|
double nrows;
|
|
|
|
/*
|
|
* Compute joinclause selectivity. Note that we are only considering
|
|
* clauses that become restriction clauses at this join level; we are not
|
|
* double-counting them because they were not considered in estimating the
|
|
* sizes of the component rels.
|
|
*
|
|
* First, see whether any of the joinclauses can be matched to known FK
|
|
* constraints. If so, drop those clauses from the restrictlist, and
|
|
* instead estimate their selectivity using FK semantics. (We do this
|
|
* without regard to whether said clauses are local or "pushed down".
|
|
* Probably, an FK-matching clause could never be seen as pushed down at
|
|
* an outer join, since it would be strict and hence would be grounds for
|
|
* join strength reduction.) fkselec gets the net selectivity for
|
|
* FK-matching clauses, or 1.0 if there are none.
|
|
*/
|
|
fkselec = get_foreign_key_join_selectivity(root,
|
|
outer_rel->relids,
|
|
inner_rel->relids,
|
|
sjinfo,
|
|
&restrictlist);
|
|
|
|
/*
|
|
* For an outer join, we have to distinguish the selectivity of the join's
|
|
* own clauses (JOIN/ON conditions) from any clauses that were "pushed
|
|
* down". For inner joins we just count them all as joinclauses.
|
|
*/
|
|
if (IS_OUTER_JOIN(jointype))
|
|
{
|
|
List *joinquals = NIL;
|
|
List *pushedquals = NIL;
|
|
ListCell *l;
|
|
|
|
/* Grovel through the clauses to separate into two lists */
|
|
foreach(l, restrictlist)
|
|
{
|
|
RestrictInfo *rinfo = lfirst_node(RestrictInfo, l);
|
|
|
|
if (RINFO_IS_PUSHED_DOWN(rinfo, joinrel->relids))
|
|
pushedquals = lappend(pushedquals, rinfo);
|
|
else
|
|
joinquals = lappend(joinquals, rinfo);
|
|
}
|
|
|
|
/* Get the separate selectivities */
|
|
jselec = clauselist_selectivity(root,
|
|
joinquals,
|
|
0,
|
|
jointype,
|
|
sjinfo);
|
|
pselec = clauselist_selectivity(root,
|
|
pushedquals,
|
|
0,
|
|
jointype,
|
|
sjinfo);
|
|
|
|
/* Avoid leaking a lot of ListCells */
|
|
list_free(joinquals);
|
|
list_free(pushedquals);
|
|
}
|
|
else
|
|
{
|
|
jselec = clauselist_selectivity(root,
|
|
restrictlist,
|
|
0,
|
|
jointype,
|
|
sjinfo);
|
|
pselec = 0.0; /* not used, keep compiler quiet */
|
|
}
|
|
|
|
/*
|
|
* Basically, we multiply size of Cartesian product by selectivity.
|
|
*
|
|
* If we are doing an outer join, take that into account: the joinqual
|
|
* selectivity has to be clamped using the knowledge that the output must
|
|
* be at least as large as the non-nullable input. However, any
|
|
* pushed-down quals are applied after the outer join, so their
|
|
* selectivity applies fully.
|
|
*
|
|
* For JOIN_SEMI and JOIN_ANTI, the selectivity is defined as the fraction
|
|
* of LHS rows that have matches, and we apply that straightforwardly.
|
|
*/
|
|
switch (jointype)
|
|
{
|
|
case JOIN_INNER:
|
|
nrows = outer_rows * inner_rows * fkselec * jselec;
|
|
/* pselec not used */
|
|
break;
|
|
case JOIN_LEFT:
|
|
nrows = outer_rows * inner_rows * fkselec * jselec;
|
|
if (nrows < outer_rows)
|
|
nrows = outer_rows;
|
|
nrows *= pselec;
|
|
break;
|
|
case JOIN_FULL:
|
|
nrows = outer_rows * inner_rows * fkselec * jselec;
|
|
if (nrows < outer_rows)
|
|
nrows = outer_rows;
|
|
if (nrows < inner_rows)
|
|
nrows = inner_rows;
|
|
nrows *= pselec;
|
|
break;
|
|
case JOIN_SEMI:
|
|
nrows = outer_rows * fkselec * jselec;
|
|
/* pselec not used */
|
|
break;
|
|
case JOIN_ANTI:
|
|
nrows = outer_rows * (1.0 - fkselec * jselec);
|
|
nrows *= pselec;
|
|
break;
|
|
default:
|
|
/* other values not expected here */
|
|
elog(ERROR, "unrecognized join type: %d", (int) jointype);
|
|
nrows = 0; /* keep compiler quiet */
|
|
break;
|
|
}
|
|
|
|
return clamp_row_est(nrows);
|
|
}
|
|
|
|
/*
|
|
* get_foreign_key_join_selectivity
|
|
* Estimate join selectivity for foreign-key-related clauses.
|
|
*
|
|
* Remove any clauses that can be matched to FK constraints from *restrictlist,
|
|
* and return a substitute estimate of their selectivity. 1.0 is returned
|
|
* when there are no such clauses.
|
|
*
|
|
* The reason for treating such clauses specially is that we can get better
|
|
* estimates this way than by relying on clauselist_selectivity(), especially
|
|
* for multi-column FKs where that function's assumption that the clauses are
|
|
* independent falls down badly. But even with single-column FKs, we may be
|
|
* able to get a better answer when the pg_statistic stats are missing or out
|
|
* of date.
|
|
*/
|
|
static Selectivity
|
|
get_foreign_key_join_selectivity(PlannerInfo *root,
|
|
Relids outer_relids,
|
|
Relids inner_relids,
|
|
SpecialJoinInfo *sjinfo,
|
|
List **restrictlist)
|
|
{
|
|
Selectivity fkselec = 1.0;
|
|
JoinType jointype = sjinfo->jointype;
|
|
List *worklist = *restrictlist;
|
|
ListCell *lc;
|
|
|
|
/* Consider each FK constraint that is known to match the query */
|
|
foreach(lc, root->fkey_list)
|
|
{
|
|
ForeignKeyOptInfo *fkinfo = (ForeignKeyOptInfo *) lfirst(lc);
|
|
bool ref_is_outer;
|
|
List *removedlist;
|
|
ListCell *cell;
|
|
|
|
/*
|
|
* This FK is not relevant unless it connects a baserel on one side of
|
|
* this join to a baserel on the other side.
|
|
*/
|
|
if (bms_is_member(fkinfo->con_relid, outer_relids) &&
|
|
bms_is_member(fkinfo->ref_relid, inner_relids))
|
|
ref_is_outer = false;
|
|
else if (bms_is_member(fkinfo->ref_relid, outer_relids) &&
|
|
bms_is_member(fkinfo->con_relid, inner_relids))
|
|
ref_is_outer = true;
|
|
else
|
|
continue;
|
|
|
|
/*
|
|
* If we're dealing with a semi/anti join, and the FK's referenced
|
|
* relation is on the outside, then knowledge of the FK doesn't help
|
|
* us figure out what we need to know (which is the fraction of outer
|
|
* rows that have matches). On the other hand, if the referenced rel
|
|
* is on the inside, then all outer rows must have matches in the
|
|
* referenced table (ignoring nulls). But any restriction or join
|
|
* clauses that filter that table will reduce the fraction of matches.
|
|
* We can account for restriction clauses, but it's too hard to guess
|
|
* how many table rows would get through a join that's inside the RHS.
|
|
* Hence, if either case applies, punt and ignore the FK.
|
|
*/
|
|
if ((jointype == JOIN_SEMI || jointype == JOIN_ANTI) &&
|
|
(ref_is_outer || bms_membership(inner_relids) != BMS_SINGLETON))
|
|
continue;
|
|
|
|
/*
|
|
* Modify the restrictlist by removing clauses that match the FK (and
|
|
* putting them into removedlist instead). It seems unsafe to modify
|
|
* the originally-passed List structure, so we make a shallow copy the
|
|
* first time through.
|
|
*/
|
|
if (worklist == *restrictlist)
|
|
worklist = list_copy(worklist);
|
|
|
|
removedlist = NIL;
|
|
foreach(cell, worklist)
|
|
{
|
|
RestrictInfo *rinfo = (RestrictInfo *) lfirst(cell);
|
|
bool remove_it = false;
|
|
int i;
|
|
|
|
/* Drop this clause if it matches any column of the FK */
|
|
for (i = 0; i < fkinfo->nkeys; i++)
|
|
{
|
|
if (rinfo->parent_ec)
|
|
{
|
|
/*
|
|
* EC-derived clauses can only match by EC. It is okay to
|
|
* consider any clause derived from the same EC as
|
|
* matching the FK: even if equivclass.c chose to generate
|
|
* a clause equating some other pair of Vars, it could
|
|
* have generated one equating the FK's Vars. So for
|
|
* purposes of estimation, we can act as though it did so.
|
|
*
|
|
* Note: checking parent_ec is a bit of a cheat because
|
|
* there are EC-derived clauses that don't have parent_ec
|
|
* set; but such clauses must compare expressions that
|
|
* aren't just Vars, so they cannot match the FK anyway.
|
|
*/
|
|
if (fkinfo->eclass[i] == rinfo->parent_ec)
|
|
{
|
|
remove_it = true;
|
|
break;
|
|
}
|
|
}
|
|
else
|
|
{
|
|
/*
|
|
* Otherwise, see if rinfo was previously matched to FK as
|
|
* a "loose" clause.
|
|
*/
|
|
if (list_member_ptr(fkinfo->rinfos[i], rinfo))
|
|
{
|
|
remove_it = true;
|
|
break;
|
|
}
|
|
}
|
|
}
|
|
if (remove_it)
|
|
{
|
|
worklist = foreach_delete_current(worklist, cell);
|
|
removedlist = lappend(removedlist, rinfo);
|
|
}
|
|
}
|
|
|
|
/*
|
|
* If we failed to remove all the matching clauses we expected to
|
|
* find, chicken out and ignore this FK; applying its selectivity
|
|
* might result in double-counting. Put any clauses we did manage to
|
|
* remove back into the worklist.
|
|
*
|
|
* Since the matching clauses are known not outerjoin-delayed, they
|
|
* would normally have appeared in the initial joinclause list. If we
|
|
* didn't find them, there are two possibilities:
|
|
*
|
|
* 1. If the FK match is based on an EC that is ec_has_const, it won't
|
|
* have generated any join clauses at all. We discount such ECs while
|
|
* checking to see if we have "all" the clauses. (Below, we'll adjust
|
|
* the selectivity estimate for this case.)
|
|
*
|
|
* 2. The clauses were matched to some other FK in a previous
|
|
* iteration of this loop, and thus removed from worklist. (A likely
|
|
* case is that two FKs are matched to the same EC; there will be only
|
|
* one EC-derived clause in the initial list, so the first FK will
|
|
* consume it.) Applying both FKs' selectivity independently risks
|
|
* underestimating the join size; in particular, this would undo one
|
|
* of the main things that ECs were invented for, namely to avoid
|
|
* double-counting the selectivity of redundant equality conditions.
|
|
* Later we might think of a reasonable way to combine the estimates,
|
|
* but for now, just punt, since this is a fairly uncommon situation.
|
|
*/
|
|
if (removedlist == NIL ||
|
|
list_length(removedlist) !=
|
|
(fkinfo->nmatched_ec - fkinfo->nconst_ec + fkinfo->nmatched_ri))
|
|
{
|
|
worklist = list_concat(worklist, removedlist);
|
|
continue;
|
|
}
|
|
|
|
/*
|
|
* Finally we get to the payoff: estimate selectivity using the
|
|
* knowledge that each referencing row will match exactly one row in
|
|
* the referenced table.
|
|
*
|
|
* XXX that's not true in the presence of nulls in the referencing
|
|
* column(s), so in principle we should derate the estimate for those.
|
|
* However (1) if there are any strict restriction clauses for the
|
|
* referencing column(s) elsewhere in the query, derating here would
|
|
* be double-counting the null fraction, and (2) it's not very clear
|
|
* how to combine null fractions for multiple referencing columns. So
|
|
* we do nothing for now about correcting for nulls.
|
|
*
|
|
* XXX another point here is that if either side of an FK constraint
|
|
* is an inheritance parent, we estimate as though the constraint
|
|
* covers all its children as well. This is not an unreasonable
|
|
* assumption for a referencing table, ie the user probably applied
|
|
* identical constraints to all child tables (though perhaps we ought
|
|
* to check that). But it's not possible to have done that for a
|
|
* referenced table. Fortunately, precisely because that doesn't
|
|
* work, it is uncommon in practice to have an FK referencing a parent
|
|
* table. So, at least for now, disregard inheritance here.
|
|
*/
|
|
if (jointype == JOIN_SEMI || jointype == JOIN_ANTI)
|
|
{
|
|
/*
|
|
* For JOIN_SEMI and JOIN_ANTI, we only get here when the FK's
|
|
* referenced table is exactly the inside of the join. The join
|
|
* selectivity is defined as the fraction of LHS rows that have
|
|
* matches. The FK implies that every LHS row has a match *in the
|
|
* referenced table*; but any restriction clauses on it will
|
|
* reduce the number of matches. Hence we take the join
|
|
* selectivity as equal to the selectivity of the table's
|
|
* restriction clauses, which is rows / tuples; but we must guard
|
|
* against tuples == 0.
|
|
*/
|
|
RelOptInfo *ref_rel = find_base_rel(root, fkinfo->ref_relid);
|
|
double ref_tuples = Max(ref_rel->tuples, 1.0);
|
|
|
|
fkselec *= ref_rel->rows / ref_tuples;
|
|
}
|
|
else
|
|
{
|
|
/*
|
|
* Otherwise, selectivity is exactly 1/referenced-table-size; but
|
|
* guard against tuples == 0. Note we should use the raw table
|
|
* tuple count, not any estimate of its filtered or joined size.
|
|
*/
|
|
RelOptInfo *ref_rel = find_base_rel(root, fkinfo->ref_relid);
|
|
double ref_tuples = Max(ref_rel->tuples, 1.0);
|
|
|
|
fkselec *= 1.0 / ref_tuples;
|
|
}
|
|
|
|
/*
|
|
* If any of the FK columns participated in ec_has_const ECs, then
|
|
* equivclass.c will have generated "var = const" restrictions for
|
|
* each side of the join, thus reducing the sizes of both input
|
|
* relations. Taking the fkselec at face value would amount to
|
|
* double-counting the selectivity of the constant restriction for the
|
|
* referencing Var. Hence, look for the restriction clause(s) that
|
|
* were applied to the referencing Var(s), and divide out their
|
|
* selectivity to correct for this.
|
|
*/
|
|
if (fkinfo->nconst_ec > 0)
|
|
{
|
|
for (int i = 0; i < fkinfo->nkeys; i++)
|
|
{
|
|
EquivalenceClass *ec = fkinfo->eclass[i];
|
|
|
|
if (ec && ec->ec_has_const)
|
|
{
|
|
EquivalenceMember *em = fkinfo->fk_eclass_member[i];
|
|
RestrictInfo *rinfo = find_derived_clause_for_ec_member(ec,
|
|
em);
|
|
|
|
if (rinfo)
|
|
{
|
|
Selectivity s0;
|
|
|
|
s0 = clause_selectivity(root,
|
|
(Node *) rinfo,
|
|
0,
|
|
jointype,
|
|
sjinfo);
|
|
if (s0 > 0)
|
|
fkselec /= s0;
|
|
}
|
|
}
|
|
}
|
|
}
|
|
}
|
|
|
|
*restrictlist = worklist;
|
|
CLAMP_PROBABILITY(fkselec);
|
|
return fkselec;
|
|
}
|
|
|
|
/*
|
|
* set_subquery_size_estimates
|
|
* Set the size estimates for a base relation that is a subquery.
|
|
*
|
|
* The rel's targetlist and restrictinfo list must have been constructed
|
|
* already, and the Paths for the subquery must have been completed.
|
|
* We look at the subquery's PlannerInfo to extract data.
|
|
*
|
|
* We set the same fields as set_baserel_size_estimates.
|
|
*/
|
|
void
|
|
set_subquery_size_estimates(PlannerInfo *root, RelOptInfo *rel)
|
|
{
|
|
PlannerInfo *subroot = rel->subroot;
|
|
RelOptInfo *sub_final_rel;
|
|
ListCell *lc;
|
|
|
|
/* Should only be applied to base relations that are subqueries */
|
|
Assert(rel->relid > 0);
|
|
Assert(planner_rt_fetch(rel->relid, root)->rtekind == RTE_SUBQUERY);
|
|
|
|
/*
|
|
* Copy raw number of output rows from subquery. All of its paths should
|
|
* have the same output rowcount, so just look at cheapest-total.
|
|
*/
|
|
sub_final_rel = fetch_upper_rel(subroot, UPPERREL_FINAL, NULL);
|
|
rel->tuples = sub_final_rel->cheapest_total_path->rows;
|
|
|
|
/*
|
|
* Compute per-output-column width estimates by examining the subquery's
|
|
* targetlist. For any output that is a plain Var, get the width estimate
|
|
* that was made while planning the subquery. Otherwise, we leave it to
|
|
* set_rel_width to fill in a datatype-based default estimate.
|
|
*/
|
|
foreach(lc, subroot->parse->targetList)
|
|
{
|
|
TargetEntry *te = lfirst_node(TargetEntry, lc);
|
|
Node *texpr = (Node *) te->expr;
|
|
int32 item_width = 0;
|
|
|
|
/* junk columns aren't visible to upper query */
|
|
if (te->resjunk)
|
|
continue;
|
|
|
|
/*
|
|
* The subquery could be an expansion of a view that's had columns
|
|
* added to it since the current query was parsed, so that there are
|
|
* non-junk tlist columns in it that don't correspond to any column
|
|
* visible at our query level. Ignore such columns.
|
|
*/
|
|
if (te->resno < rel->min_attr || te->resno > rel->max_attr)
|
|
continue;
|
|
|
|
/*
|
|
* XXX This currently doesn't work for subqueries containing set
|
|
* operations, because the Vars in their tlists are bogus references
|
|
* to the first leaf subquery, which wouldn't give the right answer
|
|
* even if we could still get to its PlannerInfo.
|
|
*
|
|
* Also, the subquery could be an appendrel for which all branches are
|
|
* known empty due to constraint exclusion, in which case
|
|
* set_append_rel_pathlist will have left the attr_widths set to zero.
|
|
*
|
|
* In either case, we just leave the width estimate zero until
|
|
* set_rel_width fixes it.
|
|
*/
|
|
if (IsA(texpr, Var) &&
|
|
subroot->parse->setOperations == NULL)
|
|
{
|
|
Var *var = (Var *) texpr;
|
|
RelOptInfo *subrel = find_base_rel(subroot, var->varno);
|
|
|
|
item_width = subrel->attr_widths[var->varattno - subrel->min_attr];
|
|
}
|
|
rel->attr_widths[te->resno - rel->min_attr] = item_width;
|
|
}
|
|
|
|
/* Now estimate number of output rows, etc */
|
|
set_baserel_size_estimates(root, rel);
|
|
}
|
|
|
|
/*
|
|
* set_function_size_estimates
|
|
* Set the size estimates for a base relation that is a function call.
|
|
*
|
|
* The rel's targetlist and restrictinfo list must have been constructed
|
|
* already.
|
|
*
|
|
* We set the same fields as set_baserel_size_estimates.
|
|
*/
|
|
void
|
|
set_function_size_estimates(PlannerInfo *root, RelOptInfo *rel)
|
|
{
|
|
RangeTblEntry *rte;
|
|
ListCell *lc;
|
|
|
|
/* Should only be applied to base relations that are functions */
|
|
Assert(rel->relid > 0);
|
|
rte = planner_rt_fetch(rel->relid, root);
|
|
Assert(rte->rtekind == RTE_FUNCTION);
|
|
|
|
/*
|
|
* Estimate number of rows the functions will return. The rowcount of the
|
|
* node is that of the largest function result.
|
|
*/
|
|
rel->tuples = 0;
|
|
foreach(lc, rte->functions)
|
|
{
|
|
RangeTblFunction *rtfunc = (RangeTblFunction *) lfirst(lc);
|
|
double ntup = expression_returns_set_rows(root, rtfunc->funcexpr);
|
|
|
|
if (ntup > rel->tuples)
|
|
rel->tuples = ntup;
|
|
}
|
|
|
|
/* Now estimate number of output rows, etc */
|
|
set_baserel_size_estimates(root, rel);
|
|
}
|
|
|
|
/*
|
|
* set_function_size_estimates
|
|
* Set the size estimates for a base relation that is a function call.
|
|
*
|
|
* The rel's targetlist and restrictinfo list must have been constructed
|
|
* already.
|
|
*
|
|
* We set the same fields as set_tablefunc_size_estimates.
|
|
*/
|
|
void
|
|
set_tablefunc_size_estimates(PlannerInfo *root, RelOptInfo *rel)
|
|
{
|
|
/* Should only be applied to base relations that are functions */
|
|
Assert(rel->relid > 0);
|
|
Assert(planner_rt_fetch(rel->relid, root)->rtekind == RTE_TABLEFUNC);
|
|
|
|
rel->tuples = 100;
|
|
|
|
/* Now estimate number of output rows, etc */
|
|
set_baserel_size_estimates(root, rel);
|
|
}
|
|
|
|
/*
|
|
* set_values_size_estimates
|
|
* Set the size estimates for a base relation that is a values list.
|
|
*
|
|
* The rel's targetlist and restrictinfo list must have been constructed
|
|
* already.
|
|
*
|
|
* We set the same fields as set_baserel_size_estimates.
|
|
*/
|
|
void
|
|
set_values_size_estimates(PlannerInfo *root, RelOptInfo *rel)
|
|
{
|
|
RangeTblEntry *rte;
|
|
|
|
/* Should only be applied to base relations that are values lists */
|
|
Assert(rel->relid > 0);
|
|
rte = planner_rt_fetch(rel->relid, root);
|
|
Assert(rte->rtekind == RTE_VALUES);
|
|
|
|
/*
|
|
* Estimate number of rows the values list will return. We know this
|
|
* precisely based on the list length (well, barring set-returning
|
|
* functions in list items, but that's a refinement not catered for
|
|
* anywhere else either).
|
|
*/
|
|
rel->tuples = list_length(rte->values_lists);
|
|
|
|
/* Now estimate number of output rows, etc */
|
|
set_baserel_size_estimates(root, rel);
|
|
}
|
|
|
|
/*
|
|
* set_cte_size_estimates
|
|
* Set the size estimates for a base relation that is a CTE reference.
|
|
*
|
|
* The rel's targetlist and restrictinfo list must have been constructed
|
|
* already, and we need an estimate of the number of rows returned by the CTE
|
|
* (if a regular CTE) or the non-recursive term (if a self-reference).
|
|
*
|
|
* We set the same fields as set_baserel_size_estimates.
|
|
*/
|
|
void
|
|
set_cte_size_estimates(PlannerInfo *root, RelOptInfo *rel, double cte_rows)
|
|
{
|
|
RangeTblEntry *rte;
|
|
|
|
/* Should only be applied to base relations that are CTE references */
|
|
Assert(rel->relid > 0);
|
|
rte = planner_rt_fetch(rel->relid, root);
|
|
Assert(rte->rtekind == RTE_CTE);
|
|
|
|
if (rte->self_reference)
|
|
{
|
|
/*
|
|
* In a self-reference, we assume the average worktable size is a
|
|
* multiple of the nonrecursive term's size. The best multiplier will
|
|
* vary depending on query "fan-out", so make its value adjustable.
|
|
*/
|
|
rel->tuples = clamp_row_est(recursive_worktable_factor * cte_rows);
|
|
}
|
|
else
|
|
{
|
|
/* Otherwise just believe the CTE's rowcount estimate */
|
|
rel->tuples = cte_rows;
|
|
}
|
|
|
|
/* Now estimate number of output rows, etc */
|
|
set_baserel_size_estimates(root, rel);
|
|
}
|
|
|
|
/*
|
|
* set_namedtuplestore_size_estimates
|
|
* Set the size estimates for a base relation that is a tuplestore reference.
|
|
*
|
|
* The rel's targetlist and restrictinfo list must have been constructed
|
|
* already.
|
|
*
|
|
* We set the same fields as set_baserel_size_estimates.
|
|
*/
|
|
void
|
|
set_namedtuplestore_size_estimates(PlannerInfo *root, RelOptInfo *rel)
|
|
{
|
|
RangeTblEntry *rte;
|
|
|
|
/* Should only be applied to base relations that are tuplestore references */
|
|
Assert(rel->relid > 0);
|
|
rte = planner_rt_fetch(rel->relid, root);
|
|
Assert(rte->rtekind == RTE_NAMEDTUPLESTORE);
|
|
|
|
/*
|
|
* Use the estimate provided by the code which is generating the named
|
|
* tuplestore. In some cases, the actual number might be available; in
|
|
* others the same plan will be re-used, so a "typical" value might be
|
|
* estimated and used.
|
|
*/
|
|
rel->tuples = rte->enrtuples;
|
|
if (rel->tuples < 0)
|
|
rel->tuples = 1000;
|
|
|
|
/* Now estimate number of output rows, etc */
|
|
set_baserel_size_estimates(root, rel);
|
|
}
|
|
|
|
/*
|
|
* set_result_size_estimates
|
|
* Set the size estimates for an RTE_RESULT base relation
|
|
*
|
|
* The rel's targetlist and restrictinfo list must have been constructed
|
|
* already.
|
|
*
|
|
* We set the same fields as set_baserel_size_estimates.
|
|
*/
|
|
void
|
|
set_result_size_estimates(PlannerInfo *root, RelOptInfo *rel)
|
|
{
|
|
/* Should only be applied to RTE_RESULT base relations */
|
|
Assert(rel->relid > 0);
|
|
Assert(planner_rt_fetch(rel->relid, root)->rtekind == RTE_RESULT);
|
|
|
|
/* RTE_RESULT always generates a single row, natively */
|
|
rel->tuples = 1;
|
|
|
|
/* Now estimate number of output rows, etc */
|
|
set_baserel_size_estimates(root, rel);
|
|
}
|
|
|
|
/*
|
|
* set_foreign_size_estimates
|
|
* Set the size estimates for a base relation that is a foreign table.
|
|
*
|
|
* There is not a whole lot that we can do here; the foreign-data wrapper
|
|
* is responsible for producing useful estimates. We can do a decent job
|
|
* of estimating baserestrictcost, so we set that, and we also set up width
|
|
* using what will be purely datatype-driven estimates from the targetlist.
|
|
* There is no way to do anything sane with the rows value, so we just put
|
|
* a default estimate and hope that the wrapper can improve on it. The
|
|
* wrapper's GetForeignRelSize function will be called momentarily.
|
|
*
|
|
* The rel's targetlist and restrictinfo list must have been constructed
|
|
* already.
|
|
*/
|
|
void
|
|
set_foreign_size_estimates(PlannerInfo *root, RelOptInfo *rel)
|
|
{
|
|
/* Should only be applied to base relations */
|
|
Assert(rel->relid > 0);
|
|
|
|
rel->rows = 1000; /* entirely bogus default estimate */
|
|
|
|
cost_qual_eval(&rel->baserestrictcost, rel->baserestrictinfo, root);
|
|
|
|
set_rel_width(root, rel);
|
|
}
|
|
|
|
|
|
/*
|
|
* set_rel_width
|
|
* Set the estimated output width of a base relation.
|
|
*
|
|
* The estimated output width is the sum of the per-attribute width estimates
|
|
* for the actually-referenced columns, plus any PHVs or other expressions
|
|
* that have to be calculated at this relation. This is the amount of data
|
|
* we'd need to pass upwards in case of a sort, hash, etc.
|
|
*
|
|
* This function also sets reltarget->cost, so it's a bit misnamed now.
|
|
*
|
|
* NB: this works best on plain relations because it prefers to look at
|
|
* real Vars. For subqueries, set_subquery_size_estimates will already have
|
|
* copied up whatever per-column estimates were made within the subquery,
|
|
* and for other types of rels there isn't much we can do anyway. We fall
|
|
* back on (fairly stupid) datatype-based width estimates if we can't get
|
|
* any better number.
|
|
*
|
|
* The per-attribute width estimates are cached for possible re-use while
|
|
* building join relations or post-scan/join pathtargets.
|
|
*/
|
|
static void
|
|
set_rel_width(PlannerInfo *root, RelOptInfo *rel)
|
|
{
|
|
Oid reloid = planner_rt_fetch(rel->relid, root)->relid;
|
|
int32 tuple_width = 0;
|
|
bool have_wholerow_var = false;
|
|
ListCell *lc;
|
|
|
|
/* Vars are assumed to have cost zero, but other exprs do not */
|
|
rel->reltarget->cost.startup = 0;
|
|
rel->reltarget->cost.per_tuple = 0;
|
|
|
|
foreach(lc, rel->reltarget->exprs)
|
|
{
|
|
Node *node = (Node *) lfirst(lc);
|
|
|
|
/*
|
|
* Ordinarily, a Var in a rel's targetlist must belong to that rel;
|
|
* but there are corner cases involving LATERAL references where that
|
|
* isn't so. If the Var has the wrong varno, fall through to the
|
|
* generic case (it doesn't seem worth the trouble to be any smarter).
|
|
*/
|
|
if (IsA(node, Var) &&
|
|
((Var *) node)->varno == rel->relid)
|
|
{
|
|
Var *var = (Var *) node;
|
|
int ndx;
|
|
int32 item_width;
|
|
|
|
Assert(var->varattno >= rel->min_attr);
|
|
Assert(var->varattno <= rel->max_attr);
|
|
|
|
ndx = var->varattno - rel->min_attr;
|
|
|
|
/*
|
|
* If it's a whole-row Var, we'll deal with it below after we have
|
|
* already cached as many attr widths as possible.
|
|
*/
|
|
if (var->varattno == 0)
|
|
{
|
|
have_wholerow_var = true;
|
|
continue;
|
|
}
|
|
|
|
/*
|
|
* The width may have been cached already (especially if it's a
|
|
* subquery), so don't duplicate effort.
|
|
*/
|
|
if (rel->attr_widths[ndx] > 0)
|
|
{
|
|
tuple_width += rel->attr_widths[ndx];
|
|
continue;
|
|
}
|
|
|
|
/* Try to get column width from statistics */
|
|
if (reloid != InvalidOid && var->varattno > 0)
|
|
{
|
|
item_width = get_attavgwidth(reloid, var->varattno);
|
|
if (item_width > 0)
|
|
{
|
|
rel->attr_widths[ndx] = item_width;
|
|
tuple_width += item_width;
|
|
continue;
|
|
}
|
|
}
|
|
|
|
/*
|
|
* Not a plain relation, or can't find statistics for it. Estimate
|
|
* using just the type info.
|
|
*/
|
|
item_width = get_typavgwidth(var->vartype, var->vartypmod);
|
|
Assert(item_width > 0);
|
|
rel->attr_widths[ndx] = item_width;
|
|
tuple_width += item_width;
|
|
}
|
|
else if (IsA(node, PlaceHolderVar))
|
|
{
|
|
/*
|
|
* We will need to evaluate the PHV's contained expression while
|
|
* scanning this rel, so be sure to include it in reltarget->cost.
|
|
*/
|
|
PlaceHolderVar *phv = (PlaceHolderVar *) node;
|
|
PlaceHolderInfo *phinfo = find_placeholder_info(root, phv);
|
|
QualCost cost;
|
|
|
|
tuple_width += phinfo->ph_width;
|
|
cost_qual_eval_node(&cost, (Node *) phv->phexpr, root);
|
|
rel->reltarget->cost.startup += cost.startup;
|
|
rel->reltarget->cost.per_tuple += cost.per_tuple;
|
|
}
|
|
else
|
|
{
|
|
/*
|
|
* We could be looking at an expression pulled up from a subquery,
|
|
* or a ROW() representing a whole-row child Var, etc. Do what we
|
|
* can using the expression type information.
|
|
*/
|
|
int32 item_width;
|
|
QualCost cost;
|
|
|
|
item_width = get_typavgwidth(exprType(node), exprTypmod(node));
|
|
Assert(item_width > 0);
|
|
tuple_width += item_width;
|
|
/* Not entirely clear if we need to account for cost, but do so */
|
|
cost_qual_eval_node(&cost, node, root);
|
|
rel->reltarget->cost.startup += cost.startup;
|
|
rel->reltarget->cost.per_tuple += cost.per_tuple;
|
|
}
|
|
}
|
|
|
|
/*
|
|
* If we have a whole-row reference, estimate its width as the sum of
|
|
* per-column widths plus heap tuple header overhead.
|
|
*/
|
|
if (have_wholerow_var)
|
|
{
|
|
int32 wholerow_width = MAXALIGN(SizeofHeapTupleHeader);
|
|
|
|
if (reloid != InvalidOid)
|
|
{
|
|
/* Real relation, so estimate true tuple width */
|
|
wholerow_width += get_relation_data_width(reloid,
|
|
rel->attr_widths - rel->min_attr);
|
|
}
|
|
else
|
|
{
|
|
/* Do what we can with info for a phony rel */
|
|
AttrNumber i;
|
|
|
|
for (i = 1; i <= rel->max_attr; i++)
|
|
wholerow_width += rel->attr_widths[i - rel->min_attr];
|
|
}
|
|
|
|
rel->attr_widths[0 - rel->min_attr] = wholerow_width;
|
|
|
|
/*
|
|
* Include the whole-row Var as part of the output tuple. Yes, that
|
|
* really is what happens at runtime.
|
|
*/
|
|
tuple_width += wholerow_width;
|
|
}
|
|
|
|
Assert(tuple_width >= 0);
|
|
rel->reltarget->width = tuple_width;
|
|
}
|
|
|
|
/*
|
|
* set_pathtarget_cost_width
|
|
* Set the estimated eval cost and output width of a PathTarget tlist.
|
|
*
|
|
* As a notational convenience, returns the same PathTarget pointer passed in.
|
|
*
|
|
* Most, though not quite all, uses of this function occur after we've run
|
|
* set_rel_width() for base relations; so we can usually obtain cached width
|
|
* estimates for Vars. If we can't, fall back on datatype-based width
|
|
* estimates. Present early-planning uses of PathTargets don't need accurate
|
|
* widths badly enough to justify going to the catalogs for better data.
|
|
*/
|
|
PathTarget *
|
|
set_pathtarget_cost_width(PlannerInfo *root, PathTarget *target)
|
|
{
|
|
int32 tuple_width = 0;
|
|
ListCell *lc;
|
|
|
|
/* Vars are assumed to have cost zero, but other exprs do not */
|
|
target->cost.startup = 0;
|
|
target->cost.per_tuple = 0;
|
|
|
|
foreach(lc, target->exprs)
|
|
{
|
|
Node *node = (Node *) lfirst(lc);
|
|
|
|
tuple_width += get_expr_width(root, node);
|
|
|
|
/* For non-Vars, account for evaluation cost */
|
|
if (!IsA(node, Var))
|
|
{
|
|
QualCost cost;
|
|
|
|
cost_qual_eval_node(&cost, node, root);
|
|
target->cost.startup += cost.startup;
|
|
target->cost.per_tuple += cost.per_tuple;
|
|
}
|
|
}
|
|
|
|
Assert(tuple_width >= 0);
|
|
target->width = tuple_width;
|
|
|
|
return target;
|
|
}
|
|
|
|
/*
|
|
* get_expr_width
|
|
* Estimate the width of the given expr attempting to use the width
|
|
* cached in a Var's owning RelOptInfo, else fallback on the type's
|
|
* average width when unable to or when the given Node is not a Var.
|
|
*/
|
|
static int32
|
|
get_expr_width(PlannerInfo *root, const Node *expr)
|
|
{
|
|
int32 width;
|
|
|
|
if (IsA(expr, Var))
|
|
{
|
|
const Var *var = (const Var *) expr;
|
|
|
|
/* We should not see any upper-level Vars here */
|
|
Assert(var->varlevelsup == 0);
|
|
|
|
/* Try to get data from RelOptInfo cache */
|
|
if (!IS_SPECIAL_VARNO(var->varno) &&
|
|
var->varno < root->simple_rel_array_size)
|
|
{
|
|
RelOptInfo *rel = root->simple_rel_array[var->varno];
|
|
|
|
if (rel != NULL &&
|
|
var->varattno >= rel->min_attr &&
|
|
var->varattno <= rel->max_attr)
|
|
{
|
|
int ndx = var->varattno - rel->min_attr;
|
|
|
|
if (rel->attr_widths[ndx] > 0)
|
|
return rel->attr_widths[ndx];
|
|
}
|
|
}
|
|
|
|
/*
|
|
* No cached data available, so estimate using just the type info.
|
|
*/
|
|
width = get_typavgwidth(var->vartype, var->vartypmod);
|
|
Assert(width > 0);
|
|
|
|
return width;
|
|
}
|
|
|
|
width = get_typavgwidth(exprType(expr), exprTypmod(expr));
|
|
Assert(width > 0);
|
|
return width;
|
|
}
|
|
|
|
/*
|
|
* relation_byte_size
|
|
* Estimate the storage space in bytes for a given number of tuples
|
|
* of a given width (size in bytes).
|
|
*/
|
|
static double
|
|
relation_byte_size(double tuples, int width)
|
|
{
|
|
return tuples * (MAXALIGN(width) + MAXALIGN(SizeofHeapTupleHeader));
|
|
}
|
|
|
|
/*
|
|
* page_size
|
|
* Returns an estimate of the number of pages covered by a given
|
|
* number of tuples of a given width (size in bytes).
|
|
*/
|
|
static double
|
|
page_size(double tuples, int width)
|
|
{
|
|
return ceil(relation_byte_size(tuples, width) / BLCKSZ);
|
|
}
|
|
|
|
/*
|
|
* Estimate the fraction of the work that each worker will do given the
|
|
* number of workers budgeted for the path.
|
|
*/
|
|
static double
|
|
get_parallel_divisor(Path *path)
|
|
{
|
|
double parallel_divisor = path->parallel_workers;
|
|
|
|
/*
|
|
* Early experience with parallel query suggests that when there is only
|
|
* one worker, the leader often makes a very substantial contribution to
|
|
* executing the parallel portion of the plan, but as more workers are
|
|
* added, it does less and less, because it's busy reading tuples from the
|
|
* workers and doing whatever non-parallel post-processing is needed. By
|
|
* the time we reach 4 workers, the leader no longer makes a meaningful
|
|
* contribution. Thus, for now, estimate that the leader spends 30% of
|
|
* its time servicing each worker, and the remainder executing the
|
|
* parallel plan.
|
|
*/
|
|
if (parallel_leader_participation)
|
|
{
|
|
double leader_contribution;
|
|
|
|
leader_contribution = 1.0 - (0.3 * path->parallel_workers);
|
|
if (leader_contribution > 0)
|
|
parallel_divisor += leader_contribution;
|
|
}
|
|
|
|
return parallel_divisor;
|
|
}
|
|
|
|
/*
|
|
* compute_bitmap_pages
|
|
*
|
|
* compute number of pages fetched from heap in bitmap heap scan.
|
|
*/
|
|
double
|
|
compute_bitmap_pages(PlannerInfo *root, RelOptInfo *baserel, Path *bitmapqual,
|
|
int loop_count, Cost *cost, double *tuple)
|
|
{
|
|
Cost indexTotalCost;
|
|
Selectivity indexSelectivity;
|
|
double T;
|
|
double pages_fetched;
|
|
double tuples_fetched;
|
|
double heap_pages;
|
|
long maxentries;
|
|
|
|
/*
|
|
* Fetch total cost of obtaining the bitmap, as well as its total
|
|
* selectivity.
|
|
*/
|
|
cost_bitmap_tree_node(bitmapqual, &indexTotalCost, &indexSelectivity);
|
|
|
|
/*
|
|
* Estimate number of main-table pages fetched.
|
|
*/
|
|
tuples_fetched = clamp_row_est(indexSelectivity * baserel->tuples);
|
|
|
|
T = (baserel->pages > 1) ? (double) baserel->pages : 1.0;
|
|
|
|
/*
|
|
* For a single scan, the number of heap pages that need to be fetched is
|
|
* the same as the Mackert and Lohman formula for the case T <= b (ie, no
|
|
* re-reads needed).
|
|
*/
|
|
pages_fetched = (2.0 * T * tuples_fetched) / (2.0 * T + tuples_fetched);
|
|
|
|
/*
|
|
* Calculate the number of pages fetched from the heap. Then based on
|
|
* current work_mem estimate get the estimated maxentries in the bitmap.
|
|
* (Note that we always do this calculation based on the number of pages
|
|
* that would be fetched in a single iteration, even if loop_count > 1.
|
|
* That's correct, because only that number of entries will be stored in
|
|
* the bitmap at one time.)
|
|
*/
|
|
heap_pages = Min(pages_fetched, baserel->pages);
|
|
maxentries = tbm_calculate_entries(work_mem * 1024L);
|
|
|
|
if (loop_count > 1)
|
|
{
|
|
/*
|
|
* For repeated bitmap scans, scale up the number of tuples fetched in
|
|
* the Mackert and Lohman formula by the number of scans, so that we
|
|
* estimate the number of pages fetched by all the scans. Then
|
|
* pro-rate for one scan.
|
|
*/
|
|
pages_fetched = index_pages_fetched(tuples_fetched * loop_count,
|
|
baserel->pages,
|
|
get_indexpath_pages(bitmapqual),
|
|
root);
|
|
pages_fetched /= loop_count;
|
|
}
|
|
|
|
if (pages_fetched >= T)
|
|
pages_fetched = T;
|
|
else
|
|
pages_fetched = ceil(pages_fetched);
|
|
|
|
if (maxentries < heap_pages)
|
|
{
|
|
double exact_pages;
|
|
double lossy_pages;
|
|
|
|
/*
|
|
* Crude approximation of the number of lossy pages. Because of the
|
|
* way tbm_lossify() is coded, the number of lossy pages increases
|
|
* very sharply as soon as we run short of memory; this formula has
|
|
* that property and seems to perform adequately in testing, but it's
|
|
* possible we could do better somehow.
|
|
*/
|
|
lossy_pages = Max(0, heap_pages - maxentries / 2);
|
|
exact_pages = heap_pages - lossy_pages;
|
|
|
|
/*
|
|
* If there are lossy pages then recompute the number of tuples
|
|
* processed by the bitmap heap node. We assume here that the chance
|
|
* of a given tuple coming from an exact page is the same as the
|
|
* chance that a given page is exact. This might not be true, but
|
|
* it's not clear how we can do any better.
|
|
*/
|
|
if (lossy_pages > 0)
|
|
tuples_fetched =
|
|
clamp_row_est(indexSelectivity *
|
|
(exact_pages / heap_pages) * baserel->tuples +
|
|
(lossy_pages / heap_pages) * baserel->tuples);
|
|
}
|
|
|
|
if (cost)
|
|
*cost = indexTotalCost;
|
|
if (tuple)
|
|
*tuple = tuples_fetched;
|
|
|
|
return pages_fetched;
|
|
}
|