<!-- doc/src/sgml/perform.sgml -->

 <chapter id="performance-tips">
  <title>Performance Tips</title>

  <indexterm zone="performance-tips">
   <primary>performance</primary>
  </indexterm>

  <para>
   Query performance can be affected by many things. Some of these can
   be controlled by the user, while others are fundamental to the underlying
   design of the system.  This chapter provides some hints about understanding
   and tuning <productname>PostgreSQL</productname> performance.
  </para>

 <sect1 id="using-explain">
  <title>Using <command>EXPLAIN</command></title>

   <indexterm zone="using-explain">
    <primary>EXPLAIN</primary>
   </indexterm>

   <indexterm zone="using-explain">
    <primary>query plan</primary>
   </indexterm>

   <para>
    <productname>PostgreSQL</productname> devises a <firstterm>query
    plan</firstterm> for each query it receives.  Choosing the right
    plan to match the query structure and the properties of the data
    is absolutely critical for good performance, so the system includes
    a complex <firstterm>planner</> that tries to choose good plans.
    You can use the <xref linkend="sql-explain"> command
    to see what query plan the planner creates for any query.
    Plan-reading is an art that requires some experience to master,
    but this section attempts to cover the basics.
   </para>

   <para>
    Examples in this section are drawn from the regression test database
    after doing a <command>VACUUM ANALYZE</>, using 9.2 development sources.
    You should be able to get similar results if you try the examples
    yourself, but your estimated costs and row counts might vary slightly
    because <command>ANALYZE</>'s statistics are random samples rather
    than exact, and because costs are inherently somewhat platform-dependent.
   </para>

   <para>
    The examples use <command>EXPLAIN</>'s default <quote>text</> output
    format, which is compact and convenient for humans to read.
    If you want to feed <command>EXPLAIN</>'s output to a program for further
    analysis, you should use one of its machine-readable output formats
    (XML, JSON, or YAML) instead.
   </para>

  <sect2 id="using-explain-basics">
   <title><command>EXPLAIN</command> Basics</title>

   <para>
    The structure of a query plan is a tree of <firstterm>plan nodes</>.
    Nodes at the bottom level of the tree are scan nodes: they return raw rows
    from a table.  There are different types of scan nodes for different
    table access methods: sequential scans, index scans, and bitmap index
    scans.  There are also non-table row sources, such as <literal>VALUES</>
    clauses and set-returning functions in <literal>FROM</>, which have their
    own scan node types.
    If the query requires joining, aggregation, sorting, or other
    operations on the raw rows, then there will be additional nodes
    above the scan nodes to perform these operations.  Again,
    there is usually more than one possible way to do these operations,
    so different node types can appear here too.  The output
    of <command>EXPLAIN</command> has one line for each node in the plan
    tree, showing the basic node type plus the cost estimates that the planner
    made for the execution of that plan node.  Additional lines might appear,
    indented from the node's summary line,
    to show additional properties of the node.
    The very first line (the summary line for the topmost
    node) has the estimated total execution cost for the plan; it is this
    number that the planner seeks to minimize.
   </para>

   <para>
    Here is a trivial example, just to show what the output looks like:

<screen>
EXPLAIN SELECT * FROM tenk1;

                         QUERY PLAN
-------------------------------------------------------------
 Seq Scan on tenk1  (cost=0.00..458.00 rows=10000 width=244)
</screen>
   </para>

   <para>
    Since this query has no <literal>WHERE</> clause, it must scan all the
    rows of the table, so the planner has chosen to use a simple sequential
    scan plan.  The numbers that are quoted in parentheses are (left
    to right):

    <itemizedlist>
     <listitem>
      <para>
       Estimated start-up cost.  This is the time expended before the output
       phase can begin, e.g., time to do the sorting in a sort node.
      </para>
     </listitem>

     <listitem>
      <para>
       Estimated total cost.  This is stated on the assumption that the plan
       node is run to completion, i.e., all available rows are retrieved.
       In practice a node's parent node might stop short of reading all
       available rows (see the <literal>LIMIT</> example below).
      </para>
     </listitem>

     <listitem>
      <para>
       Estimated number of rows output by this plan node.  Again, the node
       is assumed to be run to completion.
      </para>
     </listitem>

     <listitem>
      <para>
       Estimated average width of rows output by this plan node (in bytes).
      </para>
     </listitem>
    </itemizedlist>
   </para>

   <para>
    The costs are measured in arbitrary units determined by the planner's
    cost parameters (see <xref linkend="runtime-config-query-constants">).
    Traditional practice is to measure the costs in units of disk page
    fetches; that is, <xref linkend="guc-seq-page-cost"> is conventionally
    set to <literal>1.0</> and the other cost parameters are set relative
    to that.  The examples in this section are run with the default cost
    parameters.
   </para>

   <para>
    It's important to understand that the cost of an upper-level node includes
    the cost of all its child nodes.  It's also important to realize that
    the cost only reflects things that the planner cares about.
    In particular, the cost does not consider the time spent transmitting
    result rows to the client, which could be an important
    factor in the real elapsed time; but the planner ignores it because
    it cannot change it by altering the plan.  (Every correct plan will
    output the same row set, we trust.)
   </para>

   <para>
    The <literal>rows</> value is a little tricky because it is
    not the number of rows processed or scanned by the
    plan node, but rather the number emitted by the node.  This is often
    less than the number scanned, as a result of filtering by any
    <literal>WHERE</>-clause conditions that are being applied at the node.
    Ideally the top-level rows estimate will approximate the number of rows
    actually returned, updated, or deleted by the query.
   </para>

   <para>
    Returning to our example:

<screen>
EXPLAIN SELECT * FROM tenk1;

                         QUERY PLAN
-------------------------------------------------------------
 Seq Scan on tenk1  (cost=0.00..458.00 rows=10000 width=244)
</screen>
   </para>

   <para>
    These numbers are derived very straightforwardly.  If you do:

<programlisting>
SELECT relpages, reltuples FROM pg_class WHERE relname = 'tenk1';
</programlisting>

    you will find that <classname>tenk1</classname> has 358 disk
    pages and 10000 rows.  The estimated cost is computed as (disk pages read *
    <xref linkend="guc-seq-page-cost">) + (rows scanned *
    <xref linkend="guc-cpu-tuple-cost">).  By default,
    <varname>seq_page_cost</> is 1.0 and <varname>cpu_tuple_cost</> is 0.01,
    so the estimated cost is (358 * 1.0) + (10000 * 0.01) = 458.
   </para>

   <para>
    Now let's modify the query to add a <literal>WHERE</> condition:

<screen>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 7000;

                         QUERY PLAN
------------------------------------------------------------
 Seq Scan on tenk1  (cost=0.00..483.00 rows=7001 width=244)
   Filter: (unique1 &lt; 7000)
</screen>

    Notice that the <command>EXPLAIN</> output shows the <literal>WHERE</>
    clause being applied as a <quote>filter</> condition attached to the Seq
    Scan plan node.  This means that
    the plan node checks the condition for each row it scans, and outputs
    only the ones that pass the condition.
    The estimate of output rows has been reduced because of the
    <literal>WHERE</> clause.
    However, the scan will still have to visit all 10000 rows, so the cost
    hasn't decreased; in fact it has gone up a bit (by 10000 * <xref
    linkend="guc-cpu-operator-cost">, to be exact) to reflect the extra CPU
    time spent checking the <literal>WHERE</> condition.
   </para>

   <para>
    The actual number of rows this query would select is 7000, but the <literal>rows</>
    estimate is only approximate.  If you try to duplicate this experiment,
    you will probably get a slightly different estimate; moreover, it can
    change after each <command>ANALYZE</command> command, because the
    statistics produced by <command>ANALYZE</command> are taken from a
    randomized sample of the table.
   </para>

   <para>
    Now, let's make the condition more restrictive:

<screen>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 100;

                                  QUERY PLAN
------------------------------------------------------------------------------
 Bitmap Heap Scan on tenk1  (cost=5.03..229.17 rows=101 width=244)
   Recheck Cond: (unique1 &lt; 100)
   -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..5.01 rows=101 width=0)
         Index Cond: (unique1 &lt; 100)
</screen>

    Here the planner has decided to use a two-step plan: the child plan
    node visits an index to find the locations of rows matching the index
    condition, and then the upper plan node actually fetches those rows
    from the table itself.  Fetching rows separately is much more
    expensive than reading them sequentially, but because not all the pages
    of the table have to be visited, this is still cheaper than a sequential
    scan.  (The reason for using two plan levels is that the upper plan
    node sorts the row locations identified by the index into physical order
    before reading them, to minimize the cost of separate fetches.
    The <quote>bitmap</> mentioned in the node names is the mechanism that
    does the sorting.)
   </para>

   <para>
    Now let's add another condition to the <literal>WHERE</> clause:

<screen>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 100 AND stringu1 = 'xxx';

                                  QUERY PLAN
------------------------------------------------------------------------------
 Bitmap Heap Scan on tenk1  (cost=5.01..229.40 rows=1 width=244)
   Recheck Cond: (unique1 &lt; 100)
   Filter: (stringu1 = 'xxx'::name)
   -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..5.01 rows=101 width=0)
         Index Cond: (unique1 &lt; 100)
</screen>

    The added condition <literal>stringu1 = 'xxx'</literal> reduces the
    output-rowcount estimate, but not the cost because we still have to visit
    the same set of rows.  Notice that the <literal>stringu1</> clause
    cannot be applied as an index condition, since this index is only on
    the <literal>unique1</> column.  Instead it is applied as a filter on
    the rows retrieved by the index.  Thus the cost has actually gone up
    slightly to reflect this extra checking.
   </para>

   <para>
    In some cases the planner will prefer a <quote>simple</> index scan plan:

<screen>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 = 42;

                                 QUERY PLAN
-----------------------------------------------------------------------------
 Index Scan using tenk1_unique1 on tenk1  (cost=0.00..8.27 rows=1 width=244)
   Index Cond: (unique1 = 42)
</screen>

    In this type of plan the table rows are fetched in index order, which
    makes them even more expensive to read, but there are so few that the
    extra cost of sorting the row locations is not worth it.  You'll most
    often see this plan type for queries that fetch just a single row.  It's
    also often used for queries that have an <literal>ORDER BY</> condition
    that matches the index order, because then no extra sort step is needed to
    satisfy the <literal>ORDER BY</>.
   </para>

   <para>
    If there are indexes on several columns referenced in <literal>WHERE</>,
    the planner might choose to use an AND or OR combination of the indexes:

<screen>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 100 AND unique2 &gt; 9000;

                                     QUERY PLAN
-------------------------------------------------------------------------------------
 Bitmap Heap Scan on tenk1  (cost=25.01..60.14 rows=10 width=244)
   Recheck Cond: ((unique1 &lt; 100) AND (unique2 &gt; 9000))
   -&gt;  BitmapAnd  (cost=25.01..25.01 rows=10 width=0)
         -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..5.01 rows=101 width=0)
               Index Cond: (unique1 &lt; 100)
         -&gt;  Bitmap Index Scan on tenk1_unique2  (cost=0.00..19.74 rows=999 width=0)
               Index Cond: (unique2 &gt; 9000)
</screen>

    But this requires visiting both indexes, so it's not necessarily a win
    compared to using just one index and treating the other condition as
    a filter.  If you vary the ranges involved you'll see the plan change
    accordingly.
   </para>

   <para>
    Here is an example showing the effects of <literal>LIMIT</>:

<screen>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 100 AND unique2 &gt; 9000 LIMIT 2;

                                     QUERY PLAN
-------------------------------------------------------------------------------------
 Limit  (cost=0.00..14.25 rows=2 width=244)
   -&gt;  Index Scan using tenk1_unique2 on tenk1  (cost=0.00..71.23 rows=10 width=244)
         Index Cond: (unique2 &gt; 9000)
         Filter: (unique1 &lt; 100)
</screen>
   </para>

   <para>
    This is the same query as above, but we added a <literal>LIMIT</> so that
    not all the rows need be retrieved, and the planner changed its mind about
    what to do.  Notice that the total cost and row count of the Index Scan
    node are shown as if it were run to completion.  However, the Limit node
    is expected to stop after retrieving only a fifth of those rows, so its
    total cost is only a fifth as much, and that's the actual estimated cost
    of the query.  This plan is preferred over adding a Limit node to the
    previous plan because the Limit could not avoid paying the startup cost
    of the bitmap scan, so the total cost would be something over 25 units
    with that approach.
   </para>

   <para>
    Let's try joining two tables, using the columns we have been discussing:

<screen>
EXPLAIN SELECT *
FROM tenk1 t1, tenk2 t2
WHERE t1.unique1 &lt; 10 AND t1.unique2 = t2.unique2;

                                      QUERY PLAN
--------------------------------------------------------------------------------------
 Nested Loop  (cost=4.33..118.25 rows=10 width=488)
   -&gt;  Bitmap Heap Scan on tenk1 t1  (cost=4.33..39.44 rows=10 width=244)
         Recheck Cond: (unique1 &lt; 10)
         -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..4.33 rows=10 width=0)
               Index Cond: (unique1 &lt; 10)
   -&gt;  Index Scan using tenk2_unique2 on tenk2 t2  (cost=0.00..7.87 rows=1 width=244)
         Index Cond: (unique2 = t1.unique2)
</screen>
   </para>

   <para>
    In this plan, we have a nested-loop join node with two table scans as
    inputs, or children.  The indentation of the node summary lines reflects
    the plan tree structure.  The join's first, or <quote>outer</>, child
    is a bitmap scan similar to those we saw before.  Its cost and row count
    are the same as we'd get from <literal>SELECT ... WHERE unique1 &lt; 10</>
    because we are
    applying the <literal>WHERE</> clause <literal>unique1 &lt; 10</literal>
    at that node.
    The <literal>t1.unique2 = t2.unique2</literal> clause is not relevant yet,
    so it doesn't affect the row count of the outer scan.  The nested-loop
    join node will run its second,
    or <quote>inner</> child once for each row obtained from the outer child.
    Column values from the current outer row can be plugged into the inner
    scan; here, the <literal>t1.unique2</> value from the outer row is available,
    so we get a plan and costs similar to what we saw above for a simple
    <literal>SELECT ... WHERE t2.unique2 = <replaceable>constant</></> case.
    (The estimated cost is actually a bit lower than what was seen above,
    as a result of caching that's expected to occur during the repeated
    indexscans on <literal>t2</>.)  The
    costs of the loop node are then set on the basis of the cost of the outer
    scan, plus one repetition of the inner scan for each outer row (10 * 7.87,
    here), plus a little CPU time for join processing.
   </para>

   <para>
    In this example the join's output row count is the same as the product
    of the two scans' row counts, but that's not true in all cases because
    there can be additional <literal>WHERE</> clauses that mention both tables
    and so can only be applied at the join point, not to either input scan.
    For example, if we add one more condition:

<screen>
EXPLAIN SELECT *
FROM tenk1 t1, tenk2 t2
WHERE t1.unique1 &lt; 10 AND t1.unique2 = t2.unique2 AND t1.hundred &lt; t2.hundred;

                                      QUERY PLAN
--------------------------------------------------------------------------------------
 Nested Loop  (cost=4.33..118.28 rows=3 width=488)
   Join Filter: (t1.hundred &lt; t2.hundred)
   -&gt;  Bitmap Heap Scan on tenk1 t1  (cost=4.33..39.44 rows=10 width=244)
         Recheck Cond: (unique1 &lt; 10)
         -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..4.33 rows=10 width=0)
               Index Cond: (unique1 &lt; 10)
   -&gt;  Index Scan using tenk2_unique2 on tenk2 t2  (cost=0.00..7.87 rows=1 width=244)
         Index Cond: (unique2 = t1.unique2)
</screen>

    The extra condition <literal>t1.hundred &lt; t2.hundred</literal> can't be
    tested in the <literal>tenk2_unique2</> index, so it's applied at the
    join node.  This reduces the estimated output row count of the join node,
    but does not change either input scan.
   </para>

   <para>
    When dealing with outer joins, you might see join plan nodes with both
    <quote>Join Filter</> and plain <quote>Filter</> conditions attached.
    Join Filter conditions come from the outer join's <literal>ON</> clause,
    so a row that fails the Join Filter condition could still get emitted as
    a null-extended row.  But a plain Filter condition is applied after the
    outer-join rules and so acts to remove rows unconditionally.  In an inner
    join there is no semantic difference between these types of filters.
   </para>

   <para>
    If we change the query's selectivity a bit, we might get a very different
    join plan:

<screen>
EXPLAIN SELECT *
FROM tenk1 t1, tenk2 t2
WHERE t1.unique1 &lt; 100 AND t1.unique2 = t2.unique2;

                                        QUERY PLAN
------------------------------------------------------------------------------------------
 Hash Join  (cost=230.43..713.94 rows=101 width=488)
   Hash Cond: (t2.unique2 = t1.unique2)
   -&gt;  Seq Scan on tenk2 t2  (cost=0.00..445.00 rows=10000 width=244)
   -&gt;  Hash  (cost=229.17..229.17 rows=101 width=244)
         -&gt;  Bitmap Heap Scan on tenk1 t1  (cost=5.03..229.17 rows=101 width=244)
               Recheck Cond: (unique1 &lt; 100)
               -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..5.01 rows=101 width=0)
                     Index Cond: (unique1 &lt; 100)
</screen>
   </para>

   <para>
    Here, the planner has chosen to use a hash join, in which rows of one
    table are entered into an in-memory hash table, after which the other
    table is scanned and the hash table is probed for matches to each row.
    Again note how the indentation reflects the plan structure: the bitmap
    scan on <literal>tenk1</> is the input to the Hash node, which constructs
    the hash table.  That's then returned to the Hash Join node, which reads
    rows from its outer child plan and searches the hash table for each one.
   </para>

   <para>
    Another possible type of join is a merge join, illustrated here:

<screen>
EXPLAIN SELECT *
FROM tenk1 t1, onek t2
WHERE t1.unique1 &lt; 100 AND t1.unique2 = t2.unique2;

                                        QUERY PLAN
------------------------------------------------------------------------------------------
 Merge Join  (cost=197.83..267.93 rows=10 width=488)
   Merge Cond: (t1.unique2 = t2.unique2)
   -&gt;  Index Scan using tenk1_unique2 on tenk1 t1  (cost=0.00..656.25 rows=101 width=244)
         Filter: (unique1 &lt; 100)
   -&gt;  Sort  (cost=197.83..200.33 rows=1000 width=244)
         Sort Key: t2.unique2
         -&gt;  Seq Scan on onek t2  (cost=0.00..148.00 rows=1000 width=244)
</screen>
   </para>

   <para>
    Merge join requires its input data to be sorted on the join keys.  In this
    plan the <literal>tenk1</> data is sorted by using an index scan to visit
    the rows in the correct order, but a sequential scan and sort is preferred
    for <literal>onek</>, because there are many more rows to be visited in
    that table.
    (Seqscan-and-sort frequently beats an indexscan for sorting many rows,
    because of the nonsequential disk access required by the indexscan.)
   </para>

   <para>
    One way to look at variant plans is to force the planner to disregard
    whatever strategy it thought was the cheapest, using the enable/disable
    flags described in <xref linkend="runtime-config-query-enable">.
    (This is a crude tool, but useful.  See
    also <xref linkend="explicit-joins">.)
    For example, if we're unconvinced that seqscan-and-sort is the best way to
    deal with table <literal>onek</> in the previous example, we could try

<screen>
SET enable_sort = off;

EXPLAIN SELECT *
FROM tenk1 t1, onek t2
WHERE t1.unique1 &lt; 100 AND t1.unique2 = t2.unique2;

                                        QUERY PLAN
------------------------------------------------------------------------------------------
 Merge Join  (cost=0.00..292.36 rows=10 width=488)
   Merge Cond: (t1.unique2 = t2.unique2)
   -&gt;  Index Scan using tenk1_unique2 on tenk1 t1  (cost=0.00..656.25 rows=101 width=244)
         Filter: (unique1 &lt; 100)
   -&gt;  Index Scan using onek_unique2 on onek t2  (cost=0.00..224.76 rows=1000 width=244)
</screen>

    which shows that the planner thinks that sorting <literal>onek</> by
    indexscanning is about 12% more expensive than seqscan-and-sort.
    Of course, the next question is whether it's right about that.
    We can investigate that using <command>EXPLAIN ANALYZE</>, as discussed
    below.
   </para>

  </sect2>

  <sect2 id="using-explain-analyze">
   <title><command>EXPLAIN ANALYZE</command></title>

   <para>
    It is possible to check the accuracy of the planner's estimates
    by using <command>EXPLAIN</>'s <literal>ANALYZE</> option.  With this
    option, <command>EXPLAIN</> actually executes the query, and then displays
    the true row counts and true run time accumulated within each plan node,
    along with the same estimates that a plain <command>EXPLAIN</command>
    shows.  For example, we might get a result like this:

<screen>
EXPLAIN ANALYZE SELECT *
FROM tenk1 t1, tenk2 t2
WHERE t1.unique1 &lt; 10 AND t1.unique2 = t2.unique2;

                                                           QUERY PLAN
---------------------------------------------------------------------------------------------------------------------------------
 Nested Loop  (cost=4.33..118.25 rows=10 width=488) (actual time=0.370..1.126 rows=10 loops=1)
   -&gt;  Bitmap Heap Scan on tenk1 t1  (cost=4.33..39.44 rows=10 width=244) (actual time=0.254..0.380 rows=10 loops=1)
         Recheck Cond: (unique1 &lt; 10)
         -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..4.33 rows=10 width=0) (actual time=0.164..0.164 rows=10 loops=1)
               Index Cond: (unique1 &lt; 10)
   -&gt;  Index Scan using tenk2_unique2 on tenk2 t2  (cost=0.00..7.87 rows=1 width=244) (actual time=0.041..0.048 rows=1 loops=10)
         Index Cond: (unique2 = t1.unique2)
 Total runtime: 2.414 ms
</screen>

    Note that the <quote>actual time</quote> values are in milliseconds of
    real time, whereas the <literal>cost</> estimates are expressed in
    arbitrary units; so they are unlikely to match up.
    The thing that's usually most important to look for is whether the
    estimated row counts are reasonably close to reality.  In this example
    the estimates were all dead-on, but that's quite unusual in practice.
   </para>

   <para>
    In some query plans, it is possible for a subplan node to be executed more
    than once.  For example, the inner index scan will be executed once per
    outer row in the above nested-loop plan.  In such cases, the
    <literal>loops</> value reports the
    total number of executions of the node, and the actual time and rows
    values shown are averages per-execution.  This is done to make the numbers
    comparable with the way that the cost estimates are shown.  Multiply by
    the <literal>loops</> value to get the total time actually spent in
    the node.  In the above example, we spent a total of 0.480 milliseconds
    executing the indexscans on <literal>tenk2</>.
   </para>

   <para>
    In some cases <command>EXPLAIN ANALYZE</> shows additional execution
    statistics beyond the plan node execution times and row counts.
    For example, Sort and Hash nodes provide extra information:

<screen>
EXPLAIN ANALYZE SELECT *
FROM tenk1 t1, tenk2 t2
WHERE t1.unique1 &lt; 100 AND t1.unique2 = t2.unique2 ORDER BY t1.fivethous;

                                                                 QUERY PLAN
--------------------------------------------------------------------------------------------------------------------------------------------
 Sort  (cost=717.30..717.56 rows=101 width=488) (actual time=104.950..105.327 rows=100 loops=1)
   Sort Key: t1.fivethous
   Sort Method: quicksort  Memory: 68kB
   -&gt;  Hash Join  (cost=230.43..713.94 rows=101 width=488) (actual time=3.680..102.396 rows=100 loops=1)
         Hash Cond: (t2.unique2 = t1.unique2)
         -&gt;  Seq Scan on tenk2 t2  (cost=0.00..445.00 rows=10000 width=244) (actual time=0.046..46.219 rows=10000 loops=1)
         -&gt;  Hash  (cost=229.17..229.17 rows=101 width=244) (actual time=3.184..3.184 rows=100 loops=1)
               Buckets: 1024  Batches: 1  Memory Usage: 27kB
               -&gt;  Bitmap Heap Scan on tenk1 t1  (cost=5.03..229.17 rows=101 width=244) (actual time=0.612..1.959 rows=100 loops=1)
                     Recheck Cond: (unique1 &lt; 100)
                     -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..5.01 rows=101 width=0) (actual time=0.390..0.390 rows=100 loops=1)
                           Index Cond: (unique1 &lt; 100)
 Total runtime: 107.392 ms
</screen>

    The Sort node shows the sort method used (in particular, whether the sort
    was in-memory or on-disk) and the amount of memory or disk space needed.
    The Hash node shows the number of hash buckets and batches as well as the
    peak amount of memory used for the hash table.  (If the number of batches
    exceeds one, there will also be disk space usage involved, but that is not
    shown.)
   </para>

   <para>
    Another type of extra information is the number of rows removed by a
    filter condition:

<screen>
EXPLAIN ANALYZE SELECT * FROM tenk1 WHERE ten &lt; 7;

                                                QUERY PLAN
----------------------------------------------------------------------------------------------------------
 Seq Scan on tenk1  (cost=0.00..483.00 rows=7000 width=244) (actual time=0.111..59.249 rows=7000 loops=1)
   Filter: (ten &lt; 7)
   Rows Removed by Filter: 3000
 Total runtime: 85.340 ms
</screen>

    These counts can be particularly valuable for filter conditions applied at
    join nodes.  The <quote>Rows Removed</> line only appears when at least
    one scanned row, or potential join pair in the case of a join node,
    is rejected by the filter condition.
   </para>

   <para>
    A case similar to filter conditions occurs with <quote>lossy</>
    indexscans.  For example, consider this search for polygons containing a
    specific point:

<screen>
EXPLAIN ANALYZE SELECT * FROM polygon_tbl WHERE f1 @&gt; polygon '(0.5,2.0)';

                                              QUERY PLAN
------------------------------------------------------------------------------------------------------
 Seq Scan on polygon_tbl  (cost=0.00..1.05 rows=1 width=32) (actual time=0.251..0.251 rows=0 loops=1)
   Filter: (f1 @&gt; '((0.5,2))'::polygon)
   Rows Removed by Filter: 4
 Total runtime: 0.517 ms
</screen>

    The planner thinks (quite correctly) that this sample table is too small
    to bother with an indexscan, so we have a plain sequential scan in which
    all the rows got rejected by the filter condition.  But if we force an
    indexscan to be used, we see:

<screen>
SET enable_seqscan TO off;

EXPLAIN ANALYZE SELECT * FROM polygon_tbl WHERE f1 @&gt; polygon '(0.5,2.0)';

                                                        QUERY PLAN
--------------------------------------------------------------------------------------------------------------------------
 Index Scan using gpolygonind on polygon_tbl  (cost=0.00..8.27 rows=1 width=32) (actual time=0.293..0.293 rows=0 loops=1)
   Index Cond: (f1 @&gt; '((0.5,2))'::polygon)
   Rows Removed by Index Recheck: 1
 Total runtime: 1.054 ms
</screen>

    Here we can see that the index returned one candidate row, which was
    then rejected by a recheck of the index condition.  This happens because a
    GiST index is <quote>lossy</> for polygon containment tests: it actually
    returns the rows with polygons that overlap the target, and then we have
    to do the exact containment test on those rows.
   </para>

   <para>
    <command>EXPLAIN</> has a <literal>BUFFERS</> option that can be used with
    <literal>ANALYZE</> to get even more runtime statistics:

<screen>
EXPLAIN (ANALYZE, BUFFERS) SELECT * FROM tenk1 WHERE unique1 &lt; 100 AND unique2 &gt; 9000;

                                                            QUERY PLAN
-----------------------------------------------------------------------------------------------------------------------------------
 Bitmap Heap Scan on tenk1  (cost=25.07..60.23 rows=10 width=244) (actual time=3.069..3.213 rows=10 loops=1)
   Recheck Cond: ((unique1 &lt; 100) AND (unique2 &gt; 9000))
   Buffers: shared hit=16
   -&gt;  BitmapAnd  (cost=25.07..25.07 rows=10 width=0) (actual time=2.967..2.967 rows=0 loops=1)
         Buffers: shared hit=7
         -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..5.02 rows=102 width=0) (actual time=0.732..0.732 rows=200 loops=1)
               Index Cond: (unique1 &lt; 100)
               Buffers: shared hit=2
         -&gt;  Bitmap Index Scan on tenk1_unique2  (cost=0.00..19.80 rows=1007 width=0) (actual time=2.015..2.015 rows=1009 loops=1)
               Index Cond: (unique2 &gt; 9000)
               Buffers: shared hit=5
 Total runtime: 3.917 ms
</screen>

    The numbers provided by <literal>BUFFERS</> help to identify which parts
    of the query are the most I/O-intensive.
   </para>

   <para>
    Keep in mind that because <command>EXPLAIN ANALYZE</command> actually
    runs the query, any side-effects will happen as usual, even though
    whatever results the query might output are discarded in favor of
    printing the <command>EXPLAIN</> data.  If you want to analyze a
    data-modifying query without changing your tables, you can
    roll the command back afterwards, for example:

<screen>
BEGIN;

EXPLAIN ANALYZE UPDATE tenk1 SET hundred = hundred + 1 WHERE unique1 &lt; 100;

                                                           QUERY PLAN
--------------------------------------------------------------------------------------------------------------------------------
 Update on tenk1  (cost=5.03..229.42 rows=101 width=250) (actual time=81.055..81.055 rows=0 loops=1)
   -&gt;  Bitmap Heap Scan on tenk1  (cost=5.03..229.42 rows=101 width=250) (actual time=0.766..3.396 rows=100 loops=1)
         Recheck Cond: (unique1 &lt; 100)
         -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..5.01 rows=101 width=0) (actual time=0.461..0.461 rows=100 loops=1)
               Index Cond: (unique1 &lt; 100)
 Total runtime: 81.922 ms

ROLLBACK;
</screen>
   </para>

   <para>
    As seen in this example, when the query is an <command>INSERT</>,
    <command>UPDATE</>, or <command>DELETE</> command, the actual work of
    applying the table changes is done by a top-level Insert, Update,
    or Delete plan node.  The plan nodes underneath this node perform
    the work of locating the old rows and/or computing the new data.
    So above, we see the same sort of bitmap table scan we've seen already,
    and its output is fed to an Update node that stores the updated rows.
    It's worth noting that although the data-modifying node can take a
    considerable amount of runtime (here, it's consuming the lion's share
    of the time), the planner does not currently add anything to the cost
    estimates to account for that work.  That's because the work to be done is
    the same for every correct query plan, so it doesn't affect planning
    decisions.
   </para>

   <para>
    The <literal>Total runtime</literal> shown by <command>EXPLAIN
    ANALYZE</command> includes executor start-up and shut-down time, as well
    as the time to run any triggers that are fired, but it does not include
    parsing, rewriting, or planning time.
    Time spent executing <literal>BEFORE</> triggers, if any, is included in
    the time for the related Insert, Update, or Delete node; but time
    spent executing <literal>AFTER</> triggers is not counted there because
    <literal>AFTER</> triggers are fired after completion of the whole plan.
    The total time spent in each trigger
    (either <literal>BEFORE</> or <literal>AFTER</>) is also shown separately.
    Note that deferred constraint triggers will not be executed
    until end of transaction and are thus not shown at all by
    <command>EXPLAIN ANALYZE</command>.
   </para>

  </sect2>

  <sect2 id="using-explain-caveats">
   <title>Caveats</title>

   <para>
    There are two significant ways in which run times measured by
    <command>EXPLAIN ANALYZE</command> can deviate from normal execution of
    the same query.  First, since no output rows are delivered to the client,
    network transmission costs and I/O conversion costs are not included.
    Second, the measurement overhead added by <command>EXPLAIN
    ANALYZE</command> can be significant, especially on machines with slow
    <function>gettimeofday()</> operating-system calls. You can use the
    <xref linkend="pgtesttiming"> tool to measure the overhead of timing
    on your system.
   </para>

   <para>
    <command>EXPLAIN</> results should not be extrapolated to situations
    much different from the one you are actually testing; for example,
    results on a toy-sized table cannot be assumed to apply to large tables.
    The planner's cost estimates are not linear and so it might choose
    a different plan for a larger or smaller table.  An extreme example
    is that on a table that only occupies one disk page, you'll nearly
    always get a sequential scan plan whether indexes are available or not.
    The planner realizes that it's going to take one disk page read to
    process the table in any case, so there's no value in expending additional
    page reads to look at an index.  (We saw this happening in the
    <literal>polygon_tbl</> example above.)
   </para>

   <para>
    There are cases in which the actual and estimated values won't match up
    well, but nothing is really wrong.  One such case occurs when
    plan node execution is stopped short by a <literal>LIMIT</> or similar
    effect.  For example, in the <literal>LIMIT</> query we used before,

<screen>
EXPLAIN ANALYZE SELECT * FROM tenk1 WHERE unique1 &lt; 100 AND unique2 &gt; 9000 LIMIT 2;

                                                          QUERY PLAN
-------------------------------------------------------------------------------------------------------------------------------
 Limit  (cost=0.00..14.25 rows=2 width=244) (actual time=1.652..2.293 rows=2 loops=1)
   -&gt;  Index Scan using tenk1_unique2 on tenk1  (cost=0.00..71.23 rows=10 width=244) (actual time=1.631..2.259 rows=2 loops=1)
         Index Cond: (unique2 &gt; 9000)
         Filter: (unique1 &lt; 100)
         Rows Removed by Filter: 287
 Total runtime: 2.857 ms
</screen>

    the estimated cost and rowcount for the Index Scan node are shown as
    though it were run to completion.  But in reality the Limit node stopped
    requesting rows after it got two, so the actual rowcount is only 2 and
    the runtime is less than the cost estimate would suggest.  This is not
    an estimation error, only a discrepancy in the way the estimates and true
    values are displayed.
   </para>

   <para>
    Merge joins also have measurement artifacts that can confuse the unwary.
    A merge join will stop reading one input if it's exhausted the other input
    and the next key value in the one input is greater than the last key value
    of the other input; in such a case there can be no more matches and so no
    need to scan the rest of the first input.  This results in not reading all
    of one child, with results like those mentioned for <literal>LIMIT</>.
    Also, if the outer (first) child contains rows with duplicate key values,
    the inner (second) child is backed up and rescanned for the portion of its
    rows matching that key value.  <command>EXPLAIN ANALYZE</> counts these
    repeated emissions of the same inner rows as if they were real additional
    rows.  When there are many outer duplicates, the reported actual rowcount
    for the inner child plan node can be significantly larger than the number
    of rows that are actually in the inner relation.
   </para>

   <para>
    BitmapAnd and BitmapOr nodes always report their actual rowcounts as zero,
    due to implementation limitations.
   </para>
  </sect2>

 </sect1>

 <sect1 id="planner-stats">
  <title>Statistics Used by the Planner</title>

  <indexterm zone="planner-stats">
   <primary>statistics</primary>
   <secondary>of the planner</secondary>
  </indexterm>

  <para>
   As we saw in the previous section, the query planner needs to estimate
   the number of rows retrieved by a query in order to make good choices
   of query plans.  This section provides a quick look at the statistics
   that the system uses for these estimates.
  </para>

  <para>
   One component of the statistics is the total number of entries in
   each table and index, as well as the number of disk blocks occupied
   by each table and index.  This information is kept in the table
   <link linkend="catalog-pg-class"><structname>pg_class</structname></link>,
   in the columns <structfield>reltuples</structfield> and
   <structfield>relpages</structfield>.  We can look at it with
   queries similar to this one:

<screen>
SELECT relname, relkind, reltuples, relpages
FROM pg_class
WHERE relname LIKE 'tenk1%';

       relname        | relkind | reltuples | relpages
----------------------+---------+-----------+----------
 tenk1                | r       |     10000 |      358
 tenk1_hundred        | i       |     10000 |       30
 tenk1_thous_tenthous | i       |     10000 |       30
 tenk1_unique1        | i       |     10000 |       30
 tenk1_unique2        | i       |     10000 |       30
(5 rows)
</screen>

   Here we can see that <structname>tenk1</structname> contains 10000
   rows, as do its indexes, but the indexes are (unsurprisingly) much
   smaller than the table.
  </para>

  <para>
   For efficiency reasons, <structfield>reltuples</structfield>
   and <structfield>relpages</structfield> are not updated on-the-fly,
   and so they usually contain somewhat out-of-date values.
   They are updated by <command>VACUUM</>, <command>ANALYZE</>, and a
   few DDL commands such as <command>CREATE INDEX</>.  A <command>VACUUM</>
   or <command>ANALYZE</> operation that does not scan the entire table
   (which is commonly the case) will incrementally update the
   <structfield>reltuples</structfield> count on the basis of the part
   of the table it did scan, resulting in an approximate value.
   In any case, the planner
   will scale the values it finds in <structname>pg_class</structname>
   to match the current physical table size, thus obtaining a closer
   approximation.
  </para>

  <indexterm>
   <primary>pg_statistic</primary>
  </indexterm>

  <para>
   Most queries retrieve only a fraction of the rows in a table, due
   to <literal>WHERE</> clauses that restrict the rows to be
   examined.  The planner thus needs to make an estimate of the
   <firstterm>selectivity</> of <literal>WHERE</> clauses, that is,
   the fraction of rows that match each condition in the
   <literal>WHERE</> clause.  The information used for this task is
   stored in the
   <link linkend="catalog-pg-statistic"><structname>pg_statistic</structname></link>
   system catalog.  Entries in <structname>pg_statistic</structname>
   are updated by the <command>ANALYZE</> and <command>VACUUM
   ANALYZE</> commands, and are always approximate even when freshly
   updated.
  </para>

  <indexterm>
   <primary>pg_stats</primary>
  </indexterm>

  <para>
   Rather than look at <structname>pg_statistic</structname> directly,
   it's better to look at its view
   <link linkend="view-pg-stats"><structname>pg_stats</structname></link>
   when examining the statistics manually.  <structname>pg_stats</structname>
   is designed to be more easily readable.  Furthermore,
   <structname>pg_stats</structname> is readable by all, whereas
   <structname>pg_statistic</structname> is only readable by a superuser.
   (This prevents unprivileged users from learning something about
   the contents of other people's tables from the statistics.  The
   <structname>pg_stats</structname> view is restricted to show only
   rows about tables that the current user can read.)
   For example, we might do:

<screen>
SELECT attname, inherited, n_distinct,
       array_to_string(most_common_vals, E'\n') as most_common_vals
FROM pg_stats
WHERE tablename = 'road';

 attname | inherited | n_distinct |          most_common_vals
---------+-----------+------------+------------------------------------
 name    | f         |  -0.363388 | I- 580                        Ramp+
         |           |            | I- 880                        Ramp+
         |           |            | Sp Railroad                       +
         |           |            | I- 580                            +
         |           |            | I- 680                        Ramp
 name    | t         |  -0.284859 | I- 880                        Ramp+
         |           |            | I- 580                        Ramp+
         |           |            | I- 680                        Ramp+
         |           |            | I- 580                            +
         |           |            | State Hwy 13                  Ramp
(2 rows)
</screen>

   Note that two rows are displayed for the same column, one corresponding
   to the complete inheritance hierarchy starting at the
   <literal>road</literal> table (<literal>inherited</>=<literal>t</>),
   and another one including only the <literal>road</literal> table itself
   (<literal>inherited</>=<literal>f</>).
  </para>

  <para>
   The amount of information stored in <structname>pg_statistic</structname>
   by <command>ANALYZE</>, in particular the maximum number of entries in the
   <structfield>most_common_vals</> and <structfield>histogram_bounds</>
   arrays for each column, can be set on a
   column-by-column basis using the <command>ALTER TABLE SET STATISTICS</>
   command, or globally by setting the
   <xref linkend="guc-default-statistics-target"> configuration variable.
   The default limit is presently 100 entries.  Raising the limit
   might allow more accurate planner estimates to be made, particularly for
   columns with irregular data distributions, at the price of consuming
   more space in <structname>pg_statistic</structname> and slightly more
   time to compute the estimates.  Conversely, a lower limit might be
   sufficient for columns with simple data distributions.
  </para>

  <para>
   Further details about the planner's use of statistics can be found in
   <xref linkend="planner-stats-details">.
  </para>

 </sect1>

 <sect1 id="explicit-joins">
  <title>Controlling the Planner with Explicit <literal>JOIN</> Clauses</title>

  <indexterm zone="explicit-joins">
   <primary>join</primary>
   <secondary>controlling the order</secondary>
  </indexterm>

  <para>
   It is possible
   to control the query planner to some extent by using the explicit <literal>JOIN</>
   syntax.  To see why this matters, we first need some background.
  </para>

  <para>
   In a simple join query, such as:
<programlisting>
SELECT * FROM a, b, c WHERE a.id = b.id AND b.ref = c.id;
</programlisting>
   the planner is free to join the given tables in any order.  For
   example, it could generate a query plan that joins A to B, using
   the <literal>WHERE</> condition <literal>a.id = b.id</>, and then
   joins C to this joined table, using the other <literal>WHERE</>
   condition.  Or it could join B to C and then join A to that result.
   Or it could join A to C and then join them with B &mdash; but that
   would be inefficient, since the full Cartesian product of A and C
   would have to be formed, there being no applicable condition in the
   <literal>WHERE</> clause to allow optimization of the join.  (All
   joins in the <productname>PostgreSQL</productname> executor happen
   between two input tables, so it's necessary to build up the result
   in one or another of these fashions.)  The important point is that
   these different join possibilities give semantically equivalent
   results but might have hugely different execution costs.  Therefore,
   the planner will explore all of them to try to find the most
   efficient query plan.
  </para>

  <para>
   When a query only involves two or three tables, there aren't many join
   orders to worry about.  But the number of possible join orders grows
   exponentially as the number of tables expands.  Beyond ten or so input
   tables it's no longer practical to do an exhaustive search of all the
   possibilities, and even for six or seven tables planning might take an
   annoyingly long time.  When there are too many input tables, the
   <productname>PostgreSQL</productname> planner will switch from exhaustive
   search to a <firstterm>genetic</firstterm> probabilistic search
   through a limited number of possibilities.  (The switch-over threshold is
   set by the <xref linkend="guc-geqo-threshold"> run-time
   parameter.)
   The genetic search takes less time, but it won't
   necessarily find the best possible plan.
  </para>

  <para>
   When the query involves outer joins, the planner has less freedom
   than it does for plain (inner) joins. For example, consider:
<programlisting>
SELECT * FROM a LEFT JOIN (b JOIN c ON (b.ref = c.id)) ON (a.id = b.id);
</programlisting>
   Although this query's restrictions are superficially similar to the
   previous example, the semantics are different because a row must be
   emitted for each row of A that has no matching row in the join of B and C.
   Therefore the planner has no choice of join order here: it must join
   B to C and then join A to that result.  Accordingly, this query takes
   less time to plan than the previous query.  In other cases, the planner
   might be able to determine that more than one join order is safe.
   For example, given:
<programlisting>
SELECT * FROM a LEFT JOIN b ON (a.bid = b.id) LEFT JOIN c ON (a.cid = c.id);
</programlisting>
   it is valid to join A to either B or C first.  Currently, only
   <literal>FULL JOIN</> completely constrains the join order.  Most
   practical cases involving <literal>LEFT JOIN</> or <literal>RIGHT JOIN</>
   can be rearranged to some extent.
  </para>

  <para>
   Explicit inner join syntax (<literal>INNER JOIN</>, <literal>CROSS
   JOIN</>, or unadorned <literal>JOIN</>) is semantically the same as
   listing the input relations in <literal>FROM</>, so it does not
   constrain the join order.
  </para>

  <para>
   Even though most kinds of <literal>JOIN</> don't completely constrain
   the join order, it is possible to instruct the
   <productname>PostgreSQL</productname> query planner to treat all
   <literal>JOIN</> clauses as constraining the join order anyway.
   For example, these three queries are logically equivalent:
<programlisting>
SELECT * FROM a, b, c WHERE a.id = b.id AND b.ref = c.id;
SELECT * FROM a CROSS JOIN b CROSS JOIN c WHERE a.id = b.id AND b.ref = c.id;
SELECT * FROM a JOIN (b JOIN c ON (b.ref = c.id)) ON (a.id = b.id);
</programlisting>
   But if we tell the planner to honor the <literal>JOIN</> order,
   the second and third take less time to plan than the first.  This effect
   is not worth worrying about for only three tables, but it can be a
   lifesaver with many tables.
  </para>

  <para>
   To force the planner to follow the join order laid out by explicit
   <literal>JOIN</>s,
   set the <xref linkend="guc-join-collapse-limit"> run-time parameter to 1.
   (Other possible values are discussed below.)
  </para>

  <para>
   You do not need to constrain the join order completely in order to
   cut search time, because it's OK to use <literal>JOIN</> operators
   within items of a plain <literal>FROM</> list.  For example, consider:
<programlisting>
SELECT * FROM a CROSS JOIN b, c, d, e WHERE ...;
</programlisting>
   With <varname>join_collapse_limit</> = 1, this
   forces the planner to join A to B before joining them to other tables,
   but doesn't constrain its choices otherwise.  In this example, the
   number of possible join orders is reduced by a factor of 5.
  </para>

  <para>
   Constraining the planner's search in this way is a useful technique
   both for reducing planning time and for directing the planner to a
   good query plan.  If the planner chooses a bad join order by default,
   you can force it to choose a better order via <literal>JOIN</> syntax
   &mdash; assuming that you know of a better order, that is.  Experimentation
   is recommended.
  </para>

  <para>
   A closely related issue that affects planning time is collapsing of
   subqueries into their parent query.  For example, consider:
<programlisting>
SELECT *
FROM x, y,
    (SELECT * FROM a, b, c WHERE something) AS ss
WHERE somethingelse;
</programlisting>
   This situation might arise from use of a view that contains a join;
   the view's <literal>SELECT</> rule will be inserted in place of the view
   reference, yielding a query much like the above.  Normally, the planner
   will try to collapse the subquery into the parent, yielding:
<programlisting>
SELECT * FROM x, y, a, b, c WHERE something AND somethingelse;
</programlisting>
   This usually results in a better plan than planning the subquery
   separately.  (For example, the outer <literal>WHERE</> conditions might be such that
   joining X to A first eliminates many rows of A, thus avoiding the need to
   form the full logical output of the subquery.)  But at the same time,
   we have increased the planning time; here, we have a five-way join
   problem replacing two separate three-way join problems.  Because of the
   exponential growth of the number of possibilities, this makes a big
   difference.  The planner tries to avoid getting stuck in huge join search
   problems by not collapsing a subquery if more than <varname>from_collapse_limit</>
   <literal>FROM</> items would result in the parent
   query.  You can trade off planning time against quality of plan by
   adjusting this run-time parameter up or down.
  </para>

  <para>
   <xref linkend="guc-from-collapse-limit"> and <xref
   linkend="guc-join-collapse-limit">
   are similarly named because they do almost the same thing: one controls
   when the planner will <quote>flatten out</> subqueries, and the
   other controls when it will flatten out explicit joins.  Typically
   you would either set <varname>join_collapse_limit</> equal to
   <varname>from_collapse_limit</> (so that explicit joins and subqueries
   act similarly) or set <varname>join_collapse_limit</> to 1 (if you want
   to control join order with explicit joins).  But you might set them
   differently if you are trying to fine-tune the trade-off between planning
   time and run time.
  </para>
 </sect1>

 <sect1 id="populate">
  <title>Populating a Database</title>

  <para>
   One might need to insert a large amount of data when first populating
   a database. This section contains some suggestions on how to make
   this process as efficient as possible.
  </para>

  <sect2 id="disable-autocommit">
   <title>Disable Autocommit</title>

   <indexterm>
    <primary>autocommit</primary>
    <secondary>bulk-loading data</secondary>
   </indexterm>

   <para>
    When using multiple <command>INSERT</>s, turn off autocommit and just do
    one commit at the end.  (In plain
    SQL, this means issuing <command>BEGIN</command> at the start and
    <command>COMMIT</command> at the end.  Some client libraries might
    do this behind your back, in which case you need to make sure the
    library does it when you want it done.)  If you allow each
    insertion to be committed separately,
    <productname>PostgreSQL</productname> is doing a lot of work for
    each row that is added.  An additional benefit of doing all
    insertions in one transaction is that if the insertion of one row
    were to fail then the insertion of all rows inserted up to that
    point would be rolled back, so you won't be stuck with partially
    loaded data.
   </para>
  </sect2>

  <sect2 id="populate-copy-from">
   <title>Use <command>COPY</command></title>

   <para>
    Use <xref linkend="sql-copy"> to load
    all the rows in one command, instead of using a series of
    <command>INSERT</command> commands.  The <command>COPY</command>
    command is optimized for loading large numbers of rows; it is less
    flexible than <command>INSERT</command>, but incurs significantly
    less overhead for large data loads. Since <command>COPY</command>
    is a single command, there is no need to disable autocommit if you
    use this method to populate a table.
   </para>

   <para>
    If you cannot use <command>COPY</command>, it might help to use <xref
    linkend="sql-prepare"> to create a
    prepared <command>INSERT</command> statement, and then use
    <command>EXECUTE</command> as many times as required.  This avoids
    some of the overhead of repeatedly parsing and planning
    <command>INSERT</command>. Different interfaces provide this facility
    in different ways; look for <quote>prepared statements</> in the interface
    documentation.
   </para>

   <para>
    Note that loading a large number of rows using
    <command>COPY</command> is almost always faster than using
    <command>INSERT</command>, even if <command>PREPARE</> is used and
    multiple insertions are batched into a single transaction.
   </para>

   <para>
    <command>COPY</command> is fastest when used within the same
    transaction as an earlier <command>CREATE TABLE</command> or
    <command>TRUNCATE</command> command. In such cases no WAL
    needs to be written, because in case of an error, the files
    containing the newly loaded data will be removed anyway.
    However, this consideration only applies when
    <xref linkend="guc-wal-level"> is <literal>minimal</> as all commands
    must write WAL otherwise.
   </para>

  </sect2>

  <sect2 id="populate-rm-indexes">
   <title>Remove Indexes</title>

   <para>
    If you are loading a freshly created table, the fastest method is to
    create the table, bulk load the table's data using
    <command>COPY</command>, then create any indexes needed for the
    table.  Creating an index on pre-existing data is quicker than
    updating it incrementally as each row is loaded.
   </para>

   <para>
    If you are adding large amounts of data to an existing table,
    it might be a win to drop the indexes,
    load the table, and then recreate the indexes.  Of course, the
    database performance for other users might suffer
    during the time the indexes are missing.  One should also think
    twice before dropping a unique index, since the error checking
    afforded by the unique constraint will be lost while the index is
    missing.
   </para>
  </sect2>

  <sect2 id="populate-rm-fkeys">
   <title>Remove Foreign Key Constraints</title>

   <para>
    Just as with indexes, a foreign key constraint can be checked
    <quote>in bulk</> more efficiently than row-by-row.  So it might be
    useful to drop foreign key constraints, load data, and re-create
    the constraints.  Again, there is a trade-off between data load
    speed and loss of error checking while the constraint is missing.
   </para>

   <para>
    What's more, when you load data into a table with existing foreign key
    constraints, each new row requires an entry in the server's list of
    pending trigger events (since it is the firing of a trigger that checks
    the row's foreign key constraint).  Loading many millions of rows can
    cause the trigger event queue to overflow available memory, leading to
    intolerable swapping or even outright failure of the command.  Therefore
    it may be <emphasis>necessary</>, not just desirable, to drop and re-apply
    foreign keys when loading large amounts of data.  If temporarily removing
    the constraint isn't acceptable, the only other recourse may be to split
    up the load operation into smaller transactions.
   </para>
  </sect2>

  <sect2 id="populate-work-mem">
   <title>Increase <varname>maintenance_work_mem</varname></title>

   <para>
    Temporarily increasing the <xref linkend="guc-maintenance-work-mem">
    configuration variable when loading large amounts of data can
    lead to improved performance.  This will help to speed up <command>CREATE
    INDEX</> commands and <command>ALTER TABLE ADD FOREIGN KEY</> commands.
    It won't do much for <command>COPY</> itself, so this advice is
    only useful when you are using one or both of the above techniques.
   </para>
  </sect2>

  <sect2 id="populate-checkpoint-segments">
   <title>Increase <varname>checkpoint_segments</varname></title>

   <para>
    Temporarily increasing the <xref
    linkend="guc-checkpoint-segments"> configuration variable can also
    make large data loads faster.  This is because loading a large
    amount of data into <productname>PostgreSQL</productname> will
    cause checkpoints to occur more often than the normal checkpoint
    frequency (specified by the <varname>checkpoint_timeout</varname>
    configuration variable). Whenever a checkpoint occurs, all dirty
    pages must be flushed to disk. By increasing
    <varname>checkpoint_segments</varname> temporarily during bulk
    data loads, the number of checkpoints that are required can be
    reduced.
   </para>
  </sect2>

  <sect2 id="populate-pitr">
   <title>Disable WAL Archival and Streaming Replication</title>

   <para>
    When loading large amounts of data into an installation that uses
    WAL archiving or streaming replication, it might be faster to take a
    new base backup after the load has completed than to process a large
    amount of incremental WAL data.  To prevent incremental WAL logging
    while loading, disable archiving and streaming replication, by setting
    <xref linkend="guc-wal-level"> to <literal>minimal</>,
    <xref linkend="guc-archive-mode"> to <literal>off</>, and
    <xref linkend="guc-max-wal-senders"> to zero.
    But note that changing these settings requires a server restart.
   </para>

   <para>
    Aside from avoiding the time for the archiver or WAL sender to
    process the WAL data,
    doing this will actually make certain commands faster, because they
    are designed not to write WAL at all if <varname>wal_level</varname>
    is <literal>minimal</>.  (They can guarantee crash safety more cheaply
    by doing an <function>fsync</> at the end than by writing WAL.)
    This applies to the following commands:
    <itemizedlist>
     <listitem>
      <para>
       <command>CREATE TABLE AS SELECT</command>
      </para>
     </listitem>
     <listitem>
      <para>
       <command>CREATE INDEX</command> (and variants such as
       <command>ALTER TABLE ADD PRIMARY KEY</command>)
      </para>
     </listitem>
     <listitem>
      <para>
       <command>ALTER TABLE SET TABLESPACE</command>
      </para>
     </listitem>
     <listitem>
      <para>
       <command>CLUSTER</command>
      </para>
     </listitem>
     <listitem>
      <para>
       <command>COPY FROM</command>, when the target table has been
       created or truncated earlier in the same transaction
      </para>
     </listitem>
    </itemizedlist>
   </para>
  </sect2>

  <sect2 id="populate-analyze">
   <title>Run <command>ANALYZE</command> Afterwards</title>

   <para>
    Whenever you have significantly altered the distribution of data
    within a table, running <xref linkend="sql-analyze"> is strongly recommended. This
    includes bulk loading large amounts of data into the table.  Running
    <command>ANALYZE</command> (or <command>VACUUM ANALYZE</command>)
    ensures that the planner has up-to-date statistics about the
    table.  With no statistics or obsolete statistics, the planner might
    make poor decisions during query planning, leading to poor
    performance on any tables with inaccurate or nonexistent
    statistics.  Note that if the autovacuum daemon is enabled, it might
    run <command>ANALYZE</command> automatically; see
    <xref linkend="vacuum-for-statistics">
    and <xref linkend="autovacuum"> for more information.
   </para>
  </sect2>

  <sect2 id="populate-pg-dump">
   <title>Some Notes About <application>pg_dump</></title>

   <para>
    Dump scripts generated by <application>pg_dump</> automatically apply
    several, but not all, of the above guidelines.  To reload a
    <application>pg_dump</> dump as quickly as possible, you need to
    do a few extra things manually.  (Note that these points apply while
    <emphasis>restoring</> a dump, not while <emphasis>creating</> it.
    The same points apply whether loading a text dump with
    <application>psql</> or using <application>pg_restore</> to load
    from a <application>pg_dump</> archive file.)
   </para>

   <para>
    By default, <application>pg_dump</> uses <command>COPY</>, and when
    it is generating a complete schema-and-data dump, it is careful to
    load data before creating indexes and foreign keys.  So in this case
    several guidelines are handled automatically.  What is left
    for you to do is to:
    <itemizedlist>
     <listitem>
      <para>
       Set appropriate (i.e., larger than normal) values for
       <varname>maintenance_work_mem</varname> and
       <varname>checkpoint_segments</varname>.
      </para>
     </listitem>
     <listitem>
      <para>
       If using WAL archiving or streaming replication, consider disabling
       them during the restore. To do that, set <varname>archive_mode</>
       to <literal>off</>,
       <varname>wal_level</varname> to <literal>minimal</>, and
       <varname>max_wal_senders</> to zero before loading the dump.
       Afterwards, set them back to the right values and take a fresh
       base backup.
      </para>
     </listitem>
     <listitem>
      <para>
       Consider whether the whole dump should be restored as a single
       transaction.  To do that, pass the <option>-1</> or
       <option>--single-transaction</> command-line option to
       <application>psql</> or <application>pg_restore</>. When using this
       mode, even the smallest of errors will rollback the entire restore,
       possibly discarding many hours of processing.  Depending on how
       interrelated the data is, that might seem preferable to manual cleanup,
       or not.  <command>COPY</> commands will run fastest if you use a single
       transaction and have WAL archiving turned off.
      </para>
     </listitem>
     <listitem>
      <para>
       If multiple CPUs are available in the database server, consider using
       <application>pg_restore</>'s <option>--jobs</> option.  This
       allows concurrent data loading and index creation.
      </para>
     </listitem>
     <listitem>
      <para>
       Run <command>ANALYZE</> afterwards.
      </para>
     </listitem>
    </itemizedlist>
   </para>

   <para>
    A data-only dump will still use <command>COPY</>, but it does not
    drop or recreate indexes, and it does not normally touch foreign
    keys.

     <footnote>
      <para>
       You can get the effect of disabling foreign keys by using
       the <option>--disable-triggers</> option &mdash; but realize that
       that eliminates, rather than just postpones, foreign key
       validation, and so it is possible to insert bad data if you use it.
      </para>
     </footnote>

    So when loading a data-only dump, it is up to you to drop and recreate
    indexes and foreign keys if you wish to use those techniques.
    It's still useful to increase <varname>checkpoint_segments</varname>
    while loading the data, but don't bother increasing
    <varname>maintenance_work_mem</varname>; rather, you'd do that while
    manually recreating indexes and foreign keys afterwards.
    And don't forget to <command>ANALYZE</> when you're done; see
    <xref linkend="vacuum-for-statistics">
    and <xref linkend="autovacuum"> for more information.
   </para>
  </sect2>
  </sect1>

  <sect1 id="non-durability">
   <title>Non-Durable Settings</title>

   <indexterm zone="non-durability">
    <primary>non-durable</primary>
   </indexterm>

   <para>
    Durability is a database feature that guarantees the recording of
    committed transactions even if the server crashes or loses
    power.  However, durability adds significant database overhead,
    so if your site does not require such a guarantee,
    <productname>PostgreSQL</productname> can be configured to run
    much faster.  The following are configuration changes you can make
    to improve performance in such cases.  Except as noted below, durability
    is still guaranteed in case of a crash of the database software;
    only abrupt operating system stoppage creates a risk of data loss
    or corruption when these settings are used.

    <itemizedlist>
     <listitem>
      <para>
       Place the database cluster's data directory in a memory-backed
       file system (i.e. <acronym>RAM</> disk).  This eliminates all
       database disk I/O, but limits data storage to the amount of
       available memory (and perhaps swap).
      </para>
     </listitem>

     <listitem>
      <para>
       Turn off <xref linkend="guc-fsync">;  there is no need to flush
       data to disk.
      </para>
     </listitem>

     <listitem>
      <para>
       Turn off <xref linkend="guc-full-page-writes">;  there is no need
       to guard against partial page writes.
      </para>
     </listitem>

     <listitem>
      <para>
       Increase <xref linkend="guc-checkpoint-segments"> and <xref
       linkend="guc-checkpoint-timeout"> ; this reduces the frequency
       of checkpoints, but increases the storage requirements of
       <filename>/pg_xlog</>.
      </para>
     </listitem>

     <listitem>
      <para>
       Turn off <xref linkend="guc-synchronous-commit">;  there might be no
       need to write the <acronym>WAL</acronym> to disk on every
       commit.  This setting does risk transaction loss (though not data
       corruption) in case of a crash of the <emphasis>database</> alone.
      </para>
     </listitem>
    </itemizedlist>
   </para>
  </sect1>

 </chapter>
