chapter 13: query processing

Database System Concepts, 5th Ed.

©Silberschatz, Korth and SudarshanSee www.db-book.com for conditions on re-use

Chapter 13: Query ProcessingChapter 13: Query Processing

©Silberschatz, Korth and Sudarshan13.2Database System Concepts - 5th Edition, Aug 27, 2005.

Chapter 1: Introduction Part 1: Relational databases

Chapter 2: Relational Model Chapter 3: SQL Chapter 4: Advanced SQL Chapter 5: Other Relational Languages

Part 2: Database Design Chapter 6: Database Design and the E-R Model Chapter 7: Relational Database Design Chapter 8: Application Design and Development

Part 3: Object-based databases and XML Chapter 9: Object-Based Databases Chapter 10: XML

Part 4: Data storage and querying Chapter 11: Storage and File Structure Chapter 12: Indexing and Hashing Chapter 13: Query Processing Chapter 14: Query Optimization

Part 5: Transaction management Chapter 15: Transactions Chapter 16: Concurrency control Chapter 17: Recovery System

Database System ConceptsDatabase System Concepts

Part 6: Data Mining and Information Retrieval Chapter 18: Data Analysis and Mining Chapter 19: Information Retreival

Part 7: Database system architecture Chapter 20: Database-System Architecture Chapter 21: Parallel Databases Chapter 22: Distributed Databases

Part 8: Other topics Chapter 23: Advanced Application Development Chapter 24: Advanced Data Types and New Applications Chapter 25: Advanced Transaction Processing

Part 9: Case studies Chapter 26: PostgreSQL Chapter 27: Oracle Chapter 28: IBM DB2 Chapter 29: Microsoft SQL Server

Online Appendices Appendix A: Network Model Appendix B: Hierarchical Model Appendix C: Advanced Relational Database Model


Part 4: Data storage and querying Part 4: Data storage and querying (Chapters 11 through 14).(Chapters 11 through 14).

Chapter 11: Storage and File Structure

deals with disk, file, and file-system structure.

Chapter 12: Indexing and Hashing

A variety of data-access techniques are presented including hashing and B+ tree indices.

Chapters 13: Query Processing and Chapter 14: Query Optimization

address query-evaluation algorithms, and query optimization techniques. These chapters provide an understanding of the internals of the storage and retrieval components of a database which are necessary for query processing and optimization



13.1 Overview

13.2 Measures of Query Cost

13.3 Selection Operation

13.4 Sorting

13.5 Join Operation

13.6 Other Operations

13.7 Evaluation of Expressions

13.8 Summary


Basic Steps in Query ProcessingBasic Steps in Query Processing

1. Parsing and translation

2. Optimization

3. Evaluation


Basic Steps in Query Processing (Cont.)Basic Steps in Query Processing (Cont.)

Parsing and translation First, translate the query into its internal form. This is then translated into relational algebra. Parser checks syntax, verifies relations

Optimization Enumerate all possible query-evaluation plans Compute the cost for the plans Pick up the plan having the minimum cost

Evaluation The query-execution engine takes a query-evaluation plan, executes that plan, and returns the answers to the query.


Basic Steps in QP: OptimizationBasic Steps in QP: Optimization

A relational algebra expression may have many equivalent expressions

E.g., balance2500(balance(account)) is equivalent to balance(balance2500(account))

Each relational algebra operation can be evaluated using one of several different algorithms

Correspondingly, a relational-algebra expression can be evaluated in many ways.

Query evaluation-plan.

annotated expression specifying detailed evaluation strategy

Ex1: can use an index on balance to find accounts with balance < 2500,

Ex2: can perform complete relation scan and discard accounts with balance 2500


Basic Steps: Optimization (Cont.)Basic Steps: Optimization (Cont.)

Query Optimization

Amongst all equivalent evaluation plans, choose the one with lowest cost.

Cost is estimated using statistical information from the database catalog

e.g. number of tuples in each relation, size of tuples, etc.

In this chapter we study

How to measure query costs

Algorithms for evaluating relational algebra operations

How to combine algorithms for individual operations in order to evaluate a complete expression

In Chapter 14

We study how to optimize queries, that is, how to find an evaluation plan with lowest estimated cost



13.1 Overview



13.4 Sorting

13.5 Join Operation



13.8 Summary


Measures of Query CostMeasures of Query Cost

Cost is generally measured as total elapsed time for answering query

Many factors contribute to time cost

disk accesses, CPU, or even network communication

Typically disk access is the predominant cost, and is also relatively easy to estimate. Measured by taking into account

Number of seeks X average-seek-cost

Number of blocks read X average-block-read-cost

Number of blocks written X average-block-write-cost

Cost to write a block is greater than cost to read a block

– data is read back after being written to ensure that the write was successful


Measures of Query Cost (Cont.)Measures of Query Cost (Cont.) For simplicity we just use number of block transfers from disk as the cost

measure

We ignore the difference in cost between sequential and random I/O for simplicity

We also ignore CPU costs for simplicity

Costs depends on the size of the buffer in main memory

Having more memory reduces need for disk access

Amount of real memory available to buffer depends on other concurrent OS processes, and hard to determine ahead of actual execution

We often use worst case estimates, assuming only the minimum amount of memory needed for the operation is available

Real systems take CPU cost into account, differentiate between sequential and random I/O, and take buffer size into account

We do not include cost to writing output (query result) to disk in our cost formulae



13.1 Overview



13.4 Sorting

13.5 Join Operation



13.8 Summary


Selection OperationSelection Operation

File scan

search algorithms that locate and retrieve records that fulfill a selection condition.

Algorithm A1 (linear search):

Scan each file block and test all records to see whether they satisfy the selection condition.

Cost estimate (number of disk blocks scanned) = br

br denotes number of blocks containing records from relation r

If selection is on a key attribute, cost = (br /2)

stop on finding record

Linear search can be applied regardless of

selection condition or

ordering of records in the file, or

availability of indices


Selection Operation (Cont.)Selection Operation (Cont.)

Algorithm A2 (binary search):

Applicable if selection is an equality comparison on the attribute on which file is ordered.

Assume that the blocks of a relation are stored contiguously

Cost estimate (number of disk blocks to be scanned):

log2(br) — cost of locating the first tuple by a binary search on the blocks

Plus number of blocks containing records that satisfy selection condition

– Will see how to estimate this cost in Chapter 14


Selections Using IndicesSelections Using Indices Index scan – search algorithms that use an index

selection condition must be on search-key of index. HT: height of index

A3 (primary index on candidate key, equality). Retrieve a single record that satisfies the corresponding equality condition Cost = HTi + 1

A4 (primary index on nonkey, equality) Retrieve multiple records. Records will be on consecutive blocks Cost = HTi + number of blocks containing retrieved records

A5 (equality on search-key of secondary index). Retrieve a single record if the search-key is a candidate key

Cost = HTi + 1 Retrieve multiple records if search-key is not a candidate key

Cost = HTi + number of records retrieved

– can be very expensive! each record may be on a different block

– one block access for each retrieved record


Selections Involving ComparisonsSelections Involving Comparisons

Can implement selections of the form AV (r) or A V(r) by using

a linear file scan or binary search, or by using indices in the following ways:

A6 (primary index, comparison). (Relation is sorted on A) For A V(r) use index to find first tuple v and scan relation

sequentially from there For AV (r) just scan relation sequentially till first tuple > v; do not use

index A7 (secondary index, comparison).

For A V(r) use index to find first index entry v and scan index sequentially from there, to find pointers to records.

For AV (r) just scan leaf pages of index finding pointers to records, till first entry > v

In either case, retrieve records that are pointed to

– requires an I/O for each record

– Linear file scan may be cheaper if many records are to be fetched!


Complex SelectionsComplex Selections

Conjunction: 1 2. . . n(r)

A8 (conjunctive selection using one index).

Select a combination of i and algorithms A1 through A7 that results in the

least cost fori (r).

Test other conditions on tuple after fetching it into memory buffer.

A9 (conjunctive selection using multiple-key index).

Use appropriate composite (multiple-key) index if available.

A10 (conjunctive selection by intersection of identifiers).

Requires indices with record pointers.

Use corresponding index for each condition, and take intersection of all the obtained sets of record pointers.

Then fetch records from file

If some conditions do not have appropriate indices, apply test in memory.


Complex Selections (Cont’)Complex Selections (Cont’)

Disjunction:1 2 . . . n (r).

A11 (disjunctive selection by union of identifiers).

Applicable if all conditions have available indices.

Otherwise use linear scan.

Use corresponding index for each condition, and take union of all the obtained sets of record pointers.

Then fetch records from file

Negation: (r)

Use linear scan on file

If very few records satisfy , and an index is applicable to Find satisfying records using index and fetch from file



13.1 Overview



13.4 Sorting

13.5 Join Operation



13.8 Summary


SortingSorting

We may build an index on the relation, and then use the index to read the relation in sorted order.

This may lead to one disk block access for each tuple.

For relations that fit in memory, techniques like quicksort can be used.

Pick a middle element P in an array A

Push the elements having less value than P to the left array LA

Push the elements having bigger value than P to the right array RA

Apply the same idea recursively on LA and RA

For relations that don’t fit in memory, external sort-merge is a good choice.


External Sort-MergeExternal Sort-Merge

1. Create sorted runs. Let i be 0 initially. Repeatedly do the following till the end of the relation: (a) Read M blocks of relation into memory (b) Sort the in-memory blocks (c) Write sorted data to run Ri; increment i.Let the final value of I be N

2. Merge the runs (N-way merge). We assume (for now) that N < M. 1. Use N blocks of memory to buffer input runs, and 1 block to buffer

output. Read the first block of each run into its buffer page

2. repeat Select the first record (in sort order) among all buffer pages Write the record to the output buffer. If the output buffer is full

write it to disk. Delete the record from its input buffer page.

If the buffer page becomes empty then read the next block (if any) of the run into the buffer.

3. until all input buffer pages are empty:

Let M denote memory size (in pages).


External Sort-Merge (Cont.)External Sort-Merge (Cont.) If N M, several merge passes are required.

In each pass, contiguous groups of M - 1 runs are merged.

A pass reduces the number of runs by a factor of M -1, and creates runs longer by the same factor.

E.g. If M=11, and there are 90 runs, one pass reduces the number of runs to 9, each 10 times the size of the initial runs

Repeated passes are performed till all runs have been merged into one.


External Merge Sort (Cont.)External Merge Sort (Cont.)

Cost analysis:

br : the number of blocks in R

M : the number of blocks in each run

br / M : the initial number of runs

Total number of merge passes required: logM–1(br/M)

Disk accesses for initial run creation as well as in each pass is 2br

for final pass, we don’t count write cost

– we ignore final write cost for all operations since the output of an operation may be sent to the parent operation without being written to disk

Thus total number of disk accesses for external sorting:

br ( 2 logM–1(br / M) + 1)



13.1 Overview



13.4 Sorting

13.5 Join Operation



13.8 Summary


Join OperationJoin Operation

Several different algorithms to implement joins

Nested-loop join

Block nested-loop join

Indexed nested-loop join

Merge-join

Hash-join

Choice based on cost estimate

Examples use the following information

Number of records

customer: 10,000 depositor: 5000

Number of blocks

customer: 400 depositor: 100


Nested-Loop JoinNested-Loop Join

To compute the theta join r s

for each tuple tr in r do begin

for each tuple ts in s do begin

test pair (tr,ts) to see if they satisfy the join condition

if they do, add tr • ts to the result.

endend

r is called the outer relation and s the inner relation of the join.

Requires no indices and can be used with any kind of join condition.

Expensive since it examines every pair of tuples in the two relations.


Loan number Branch name amount

L- 170 Downtown 300

L – 42 Redwood 400


L – 112 Perryridge 2300


L – 90 Downtown 800



L - 70 Perryridge 600

L - 221 Downtown 1000



Customer name

Loan number

Jones L - 170

Bahn L – 82

Kim L – 42

Lee L – 48

Jane L – 112

Smith L – 34

Hwang L – 321

Choi L – 109

Pedro L – 90

Sammy L – 112

Jun L – 31

Jung L – 62

Shin L – 99

Koh L – 70

Mark L – 221

Harry L - 116

Loan relation : N = 12, B = 3

Borrower relation: N = 16, B = 4

Nested-loop joinNested-loop join보조자료


Nested-loop joinNested-loop join


L- 170 Downtown 300












Customer name Loan number

Jones L - 170

Bahn L – 82

Kim L – 42

Lee L – 48

Jane L – 112

Smith L – 34

Hwang L – 321

Choi L – 109

Pedro L – 90

Sammy L – 112

Jun L – 31

Jung L – 62

Shin L – 99

Koh L – 70

Mark L – 221

Harry L - 116

Loan relationBorrower relation

보조자료




L- 170 Downtown 300













Jones L - 170

Bahn L – 82

Kim L – 42

Lee L – 48

Jane L – 112

Smith L – 34

Hwang L – 321

Choi L – 109

Pedro L – 90

Sammy L – 112

Jun L – 31

Jung L – 62

Shin L – 99

Koh L – 70

Mark L – 221

Harry L - 116


보조자료




L- 170 Downtown 300













Jones L - 170

Bahn L – 82

Kim L – 42

Lee L – 48

Jane L – 112

Smith L – 34

Hwang L – 321

Choi L – 109

Pedro L – 90

Sammy L – 112

Jun L – 31

Jung L – 62

Shin L – 99

Koh L – 70

Mark L – 221

Harry L - 116

Loan relation

Borrower relation

보조자료




L- 170 Downtown 300













Jones L - 170

Bahn L – 82

Kim L – 42

Lee L – 48

Jane L – 112

Smith L – 34

Hwang L – 321

Choi L – 109

Pedro L – 90

Sammy L – 112

Jun L – 31

Jung L – 62

Shin L – 99

Koh L – 70

Mark L – 221

Harry L - 116


borrower relation 을 outer relation 으로 하면 16*3 + 4 = 52 회의 disk access 가 필요

Loan relation 을 outer relation 으로 했을 경우 12*4 + 3 = 51 회의 disk access 가 필요

보조자료


Nested-Loop Join (Cont.)Nested-Loop Join (Cont.)

In the worst case, if there is enough memory only to hold one block of each relation, the estimated cost is nr bs + br disk accesses

where nr is number of record in R and bs and br are number of disk blocks in S and R

If the smaller relation fits entirely in memory, use that as the inner relation.

Reduces cost to br + bs disk accesses.

Assuming worst case memory availability cost estimate is

5000 400 + 100 = 2,000,100 disk accesses with depositor as outer relation, and

1000 100 + 400 = 1,000,400 disk accesses with customer as the outer relation.

If smaller relation (depositor) fits entirely in memory, the cost estimate will be 500 disk accesses.

Block nested-loops algorithm (next slide) is preferable.


Block Nested-Loop JoinBlock Nested-Loop Join

Variant of nested-loop join in which every block of inner relation is paired with every block of outer relation.

for each block Br of r do begin

for each block Bs of s do begin

for each tuple tr in Br do begin

for each tuple ts in Bs do begin

Check if (tr,ts) satisfy the join condition

if they do, add tr • ts to the result.

endend

endend

Won Kim’s Join Method

One chapter of ’80 PhD Thesis at Univ of Illinois at Urbana-Champaign


Block nested-loop joinBlock nested-loop join


L- 170 Downtown 300













Jones L - 170

Bahn L – 82

Kim L – 42

Lee L – 48

Jane L – 112

Smith L – 34

Hwang L – 321

Choi L – 109

Pedro L – 90

Sammy L – 112

Jun L – 31

Jung L – 62

Shin L – 99

Koh L – 70

Mark L – 221

Harry L - 116


보조자료




L- 170 Downtown 300













Jones L - 170

Bahn L – 82

Kim L – 42

Lee L – 48

Jane L – 112

Smith L – 34

Hwang L – 321

Choi L – 109

Pedro L – 90

Sammy L – 112

Jun L – 31

Jung L – 62

Shin L – 99

Koh L – 70

Mark L – 221

Harry L - 116


보조자료



Loan number

Branch name

amount

L- 170 Downtown 300













Jones L - 170

Bahn L – 82

Kim L – 42

Lee L – 48

Jane L – 112

Smith L – 34

Hwang L – 321

Choi L – 109

Pedro L – 90

Sammy L – 112

Jun L – 31

Jung L – 62

Shin L – 99

Koh L – 70

Mark L – 221

Harry L - 116


보조자료




L- 170 Downtown 300













Jones L - 170

Bahn L – 82

Kim L – 42

Lee L – 48

Jane L – 112

Smith L – 34

Hwang L – 321

Choi L – 109

Pedro L – 90

Sammy L – 112

Jun L – 31

Jung L – 62

Shin L – 99

Koh L – 70

Mark L – 221

Harry L - 116


borrower relation 을 outer relation 으로 하면 4*3 + 4 = 16 회의 disk access 가 필요

Loan relation 을 outer relation 으로 했을 경우 3*4 + 3 = 15 회의 disk access 가 필요

보조자료


Block Nested-Loop Join (Cont.)Block Nested-Loop Join (Cont.)

Worst case estimate: br bs + br block accesses. Each block in the inner relation s is read once for each block in the outer

relation (instead of once for each tuple in the outer relation

Best case: br + bs block accesses.

Improvements to nested loop and block nested loop algorithms: In block nested-loop, use M - 2 disk blocks as blocking unit for outer

relations, where M = memory size in blocks; use remaining two blocks to buffer inner relation and output

Cost = br / (M-2) bs + br If equi-join attribute forms a key of inner relation, stop inner loop on first

match Scan inner loop forward and backward alternately, to make use of the

blocks remaining in buffer (with LRU replacement) Use index on inner relation if available (next slide)


Indexed Nested-Loop JoinIndexed Nested-Loop Join

Index lookups can replace file scans if join is an equi-join or natural join and an index is available on the inner relation’s join attribute

Can construct an index just to compute a join.

For each tuple tr in the outer relation r, use the index (B+ tree) to look up tuples in s that satisfy the join condition with tuple tr.

Worst case: buffer has space for only one page of r, and for each tuple in r, we perform an index lookup on s.

Cost of the join: br + nr c Where c is the cost of traversing index and fetching all matching s tuples for

one tuple or r c can be estimated as cost of a single selection on s using the join condition.

If indices are available on join attributes of both r and s, use the relation with fewer tuples as the outer relation.


Indexed nested-loop joinIndexed nested-loop join

Borrower relation 의 16 개의 tuple 에 대해 각각 그 tuple 과 join 할 수 있는 loan relation 의 tuple 을 B+tree로 검색 (3 회의 disk access)

Tree 의 height 가 2, 그리고 , 실제의 data access 를 위한 1 회 , 따라서 , 3 회의 disk access 필요 따라서 , 총 cost 는 4 block access +16*3 block access = 43 회의 disk access 필요


L- 170 Downtown 300













Jones L - 170

Bahn L – 82

Kim L – 42

Lee L – 48

Jane L – 112

Smith L – 34

Hwang L – 321

Choi L – 109

Pedro L – 90

Sammy L – 112

Jun L – 31

Jung L – 62

Shin L – 99

Koh L – 70

Mark L – 221

Harry L - 116

B+ treeIndex

보조자료


Example of Nested-Loop Join CostsExample of Nested-Loop Join Costs

Compute depositor customer, with depositor as the outer relation.

Let customer have a primary B+-tree index on the join attribute customer-name, which contains 20 entries in each index node.

Then, since customer has 10,000 tuples, the height of the tree is 4, and one more access is needed to find the actual data

depositor has 5000 tuples

Cost of block nested-loop join

400*100 + 100 = 40,100 disk accesses assuming worst case memory (may be significantly less with more memory)

Cost of indexed nested-loop join

100 + 5000 * 5 = 25,100 disk accesses

For each depositor record, search B+ tree of customer

CPU cost likely to be less than that for block nested loops join


Merge-JoinMerge-Join

1. Sort both relations on their join attribute (if not already sorted on the join attributes).

2. Merge the sorted relations to join them

1. Join step is similar to the merge stage of the sort-merge algorithm.

2. Main difference is handling of duplicate values in join attribute — every pair with same value on join attribute must be matched

3. Detailed algorithm in book


Merge Join Merge Join on The on The borrower and loan borrower and loan RelationRelation

보조자료

Custom-name

Loan -number

Loan-number

Branch-name

amount

Smith L-11 L-11 Round Hill 900

Jackson L-14 L-14 Downtown 1500

Hayes L-15 L-15 Perryridge 1500

Adams L-16 L-16 Perryridge 1300

Jones L-17 L-17 Downtown 1000

Williams L-17

Smith L-23 L-23 Redwood 2000

Curry L-93 L-93 Mianus 500


Merge-Join (Cont.)Merge-Join (Cont.)

Can be used only for equi-joins and natural joins

Each block needs to be read only once (assuming all tuples for any given value of the join attributes fit in memory

Thus number of block accesses for merge-join is br + bs + the cost of sorting if relations are unsorted.

hybrid merge-join: If one relation is sorted, and the other has a secondary B+-tree index on the join attribute

Merge the sorted relation with the leaf entries of the B+- tree

Using the order property of the leaf entries of the B+- tree

Sort the result on the addresses of the unsorted relation’s tuples

Scan the unsorted relation in physical address order and merge with previous result, to replace addresses by the actual tuples

Sequential scan more efficient than random lookup


Hash-JoinHash-Join

Applicable for equi-joins and natural joins. A hash function h is used to partition tuples of both relations h maps JoinAttrs values to {0, 1, ..., n}, where JoinAttrs denotes the common

attributes of r and s used in the natural join.

r0, r1, . . ., rn denote partitions of r tuples

Each tuple tr r is put in partition ri where i = h (tr [JoinAttrs]).

s0,, s1. . ., sn denotes partitions of s tuples

Each tuple ts s is put in partition si, where i = h (ts [JoinAttrs]).

Note: In book, ri is denoted as Hri, si is denoted as Hsi and n is denoted as nh

r tuples in ri need only to be compared with s tuples in si

Need not be compared with s tuples in any other partition, since: an r tuple and an s tuple that satisfy the join condition will have the same

value for the join attributes..


Hash-Join (Cont.)Hash-Join (Cont.)

JOIN

JOIN

JOIN

JOIN

JOIN


Hash Join Hash Join on The on The borrower and loan borrower and loan RelationRelation

보조자료

Custom-name

Loan -number

해시집합

(R)

해시집합

(S)

Loan-number

Branch-name amount

Adams L-16 3 0 L-11 Round Hill 900

Curry L-93 6 1 L-14 Downtown 1500

Hayes L-15 2 2 L-15 Perryridge 1500

Jackson L-14 1 3 L-16 Perryridge 1300

Jones L-17 4 4 L-17 Downtown 1000

Smith L-11 0 5 L-23 Redwood 2000

Smith L-23 5 6 L-93 Mianus 500

Williams L-17 4

Custom-name

Loan -number

해시집합

(R)

해시집합

(S)

Loan-number

Branch-name amount

Smith L-11 0 0 L-11 Round Hill 900

Jackson L-14 1 1 L-14 Downtown 1500

Hayes L-15 2 2 L-15 Perryridge 1500

Adams L-16 3 3 L-16 Perryridge 1300

Jones L-17 4 4 L-17 Downtown 1000

Williams L-17 4

Smith L-23 5 5 L-23 Redwood 2000

Curry L-93 6 6 L-93 Mianus 500


Hash-Join AlgorithmHash-Join Algorithm

1. Partition the relation s using hashing function h.

When partitioning a relation, one block of memory is reserved as the output buffer for each partition.

2. Partition r similarly.

3. For each i:

(a) Load si into memory and build an in-memory hash index on it using the join attribute.

This hash index uses a different hash function than the earlier one h.

(b) Read the tuples in ri from the disk one by one.

For each tuple tr locate each matching tuple ts in si using the in-memory hash index.

Output the concatenation of their attributes.

The hash-join of r and s is computed as follows.

Relation s is called the build input and r is called the probe input.


Hash-Join Algorithm AnalysisHash-Join Algorithm Analysis

The value n and the hash function h is chosen such that each si should fit in

memory.

Typically n is chosen as bs/M * f where f is a “fudge factor”, typically

around 1.2

The probe relation partitions si need not fit in memory

Recursive partitioning required if number of partitions n is greater than number of pages M of memory.

instead of partitioning n ways, use M – 1 partitions for s

Further partition the M – 1 partitions using a different hash function

Use same partitioning method on r

Rarely required: e.g., recursive partitioning not needed for relations of 1GB or less with memory size of 2MB, with block size of 4KB.


Handling of Overflows in Hash-JoinHandling of Overflows in Hash-Join

Hash-table overflow occurs in partition si if si does not fit in memory.

Reasons could be

Many tuples in s with same value for join attributes

Bad hash function

Partitioning is said to be skewed if some partitions have significantly more tuples than some others

Overflow resolution can be done in build phase

Partition si is further partitioned using different hash function.

Partition ri must be similarly partitioned.

Overflow avoidance performs partitioning carefully to avoid overflows during build phase

E.g. partition build relation into many partitions, then combine them

Both approaches fail with large numbers of duplicates

Fallback option: use block nested loops join on overflowed partitions


Cost of Hash-JoinCost of Hash-Join

If recursive partitioning is not required: cost of hash join is 3(br + bs) +2 nh

If recursive partitioning required, number of passes required for partitioning s is logM–1(bs) – 1. This is because each final partition of s should fit in memory.

The number of partitions of probe relation r is the same as that for build relation s; the number of passes for partitioning of r is also the same as for s.

Therefore it is best to choose the smaller relation as the build relation.

Total cost estimate with recursive partitioning is:

2(br + bs logM–1(bs) – 1 + br + bs

If the entire build input can be kept in main memory, n can be set to 0 and the algorithm does not partition the relations into temporary files.

Cost estimate goes down to br + bs.


Example of Cost of Hash-JoinExample of Cost of Hash-Join

Assume that memory size is 20 blocks

bdepositor= 100 and bcustomer = 400.

depositor is to be used as build input.

Partition it into five partitions, each of size 20 blocks.

This partitioning can be done in one pass.

Similarly, partition customer into five partitions, each of size 80.

This is also done in one pass.

Therefore total cost: 3(100 + 400) = 1500 block transfers

ignores cost of writing partially filled blocks

customer depositor


Hybrid Hash–JoinHybrid Hash–Join Useful when memory sized are relatively large, and the build input is bigger than

memory. Main feature of hybrid hash join:

Keep the first partition of the build relation in memory. E.g. With memory size of 25 blocks, depositor can be partitioned into five

partitions, each of size 20 blocks. Division of memory:

The first partition occupies 20 blocks of memory 1 block is used for input, and 1 block each for buffering the other 4 partitions.

customer is similarly partitioned into five partitions each of size 80; the first is used right away for probing, instead of being written out and read

back. Cost of 3(80 + 320) + 20 +80 = 1300 block transfers for

hybrid hash join, instead of 1500 with plain hash-join.

Hybrid hash-join most useful if M >> sb


Complex JoinsComplex Joins

Join with a conjunctive condition:

r 1 2... n s

Either use nested loops/block nested loops, or

Compute the result of one of the simpler joins r i s

final result comprises those tuples in the intermediate result that satisfy the remaining conditions

1 . . . i –1 i +1 . . . n

Join with a disjunctive condition

r 1 2 ... n s

Either use nested loops/block nested loops, or

Compute as the union of the records in individual joins r i s:

(r 1 s) (r 2 s) . . . (r n s)



13.1 Overview



13.4 Sorting

13.5 Join Operation



13.8 Summary


Other Operations: Duplicate eliminationOther Operations: Duplicate elimination

Duplicate elimination can be implemented via hashing or sorting.

On sorting duplicates will come adjacent to each other, and all but one set of duplicates can be deleted.

Optimization: duplicates can be deleted during run generation as well as at intermediate merge steps in external sort-merge.

Hashing is similar – duplicates will come into the same bucket.

Projection is implemented by performing projection on each tuple followed by duplicate elimination.


Other Operations : AggregationOther Operations : Aggregation

Aggregation can be implemented in a manner similar to duplicate elimination.

Sorting or hashing can be used to bring tuples in the same group together, and then the aggregate functions can be applied on each group.

Optimization: combine tuples in the same group during run generation and intermediate merges, by computing partial aggregate values

For count, min, max, sum:

– keep aggregate values on tuples found so far in the group.

– When combining partial aggregate for count, add up the aggregates

For avg:

– keep sum and count,

– and divide sum by count at the end


Other Operations : Set OperationsOther Operations : Set Operations

Set operations (, and ):

can either use variant of merge-join after sorting, or variant of hash-join.

E.g., Set operations using hashing:

Partition both relations using the same hash function, thereby creating, r0,, r1, ..,

rn and s0 , s1, .., sn

Process each partition i as follows.

Using a different hashing function, build an in-memory hash index on ri after it is brought into memory.

r s: Add tuples in si to the hash index if they are not already in it.

At end of si add the tuples in the hash index to the result.

r s: Output tuples in si to the result if they are already there in the hash index.

r – s: For each tuple in si, if it is there in the hash index, delete it from the index.

At end of si add remaining tuples in the hash index to the result.


Other Operations : Outer JoinOther Operations : Outer Join

Outer join can be computed either as

A join followed by addition of null-padded non-participating tuples.

by modifying the join algorithms.

Modifying merge join to compute r s

In r s, non participating tuples are those in r – R(r s)

Modify merge-join to compute r s: During merging, for every tuple tr from r that do not match any tuple in s, output tr padded with nulls.

Right outer-join and full outer-join can be computed similarly.

Modifying hash join to compute r s

If r is probe relation, output non-matching r tuples padded with nulls

If r is build relation, when probing keep track of which r tuples matched s tuples. At end of si output non-matched r tuples padded with nulls



13.1 Overview



13.4 Sorting

13.5 Join Operation



13.8 Summary


Evaluation of ExpressionsEvaluation of Expressions

So far, we have seen algorithms for individual operations

Need to consider alternatives for evaluating an entire expression tree

Materialization:

generate results of an expression whose inputs are relations or are already computed,

materialize (store) it on disk. Repeat.

Pipelining:

pass on tuples to parent operations even as an operation is being executed

Alternatives for evaluating an entire expression tree may have different costs

We study above alternatives in more detail


MaterializationMaterialization Materialized evaluation:

evaluate one operation at a time, starting at the lowest-level. Use intermediate results materialized into temporary relations to evaluate next-

level operations. E.g., in figure below,

compute and store (materialize) then compute the store its join with customer, and finally compute the projections on customer-name.

)(2500 accountbalance


Materialization (Cont.)Materialization (Cont.)

Materialized evaluation is always applicable

Cost of writing results to disk and reading them back can be quite high

Our cost formulas for operations ignore cost of writing final results to disk, so

Overall cost = sum of costs of individual operations + cost of writing intermediate results to disk

Double buffering: (use two output buffers for each operation)

When one buffer is full, write it to disk while the other buffer is getting filled

Allows overlap of disk writes with computation

Reduces the total execution time


PipeliningPipelining

Pipelined evaluation :

evaluate several operations simultaneously,

passing the results of one operation on to the next.

E.g., in previous expression tree, don’t store the result of

instead, pass tuples directly to the join.

Similarly, don’t store result of join, pass tuples directly to projection.

Much cheaper than materialization: no need to store a temporary relation to disk.

Pipelining may not always be possible – e.g., sort, hash-join.

For pipelining to be effective, use evaluation algorithms that generate output tuples even as tuples are received for inputs to the operation.

Pipelines can be executed in two ways

demand driven and producer driven

)(2500 accountbalance


Pipelining (Cont.)Pipelining (Cont.) In demand driven or lazy evaluation

System repeatedly requests next tuple from top level operation

Each operation requests next tuple from children operations as required, in order to output its next tuple

In between calls, operation has to maintain “state” so it knows what to return next

Each operation is implemented as an iterator implementing the following operations

open()

– E.g. file scan: initialize file scan, store pointer to beginning of file as state

– E.g.merge join: sort relations and store pointers to beginning of sorted relations as state

next()

– E.g. for file scan: Output next tuple, and advance and store file pointer

– E.g. for merge join: continue with merge from earlier state till next output tuple is found. Save pointers as iterator state.

close()


Pipelining (Cont.)Pipelining (Cont.)

In produce-driven or eager pipelining

Operators produce tuples eagerly and pass them up to their parents

Buffer maintained between operators, child puts tuples in buffer, parent removes tuples from buffer

if buffer is full, child waits till there is space in the buffer, and then generates more tuples

System schedules operations that have space in output buffer and can process more input tuples


Evaluation Algorithms for PipeliningEvaluation Algorithms for Pipelining

Some algorithms are not able to output results even as they get input tuples

E.g. merge join, or hash join

These result in intermediate results being written to disk and then read back always

Algorithm variants are possible to generate (at least some) results on the fly, as input tuples are read in

E.g. hybrid hash join generates output tuples even as probe relation tuples in the in-memory partition (partition 0) are read in

Pipelined join technique: Hybrid hash join, modified to buffer partition 0 tuples of both relations in-memory, reading them as they become available, and output results of any matches between partition 0 tuples

When a new r0 tuple is found, match it with existing s0 tuples, output

matches, and save it in r0

Symmetrically for s0 tuples


Multiple JoinsMultiple Joins Join involving three relations: loan depositor customer

Strategy 1.

Compute depositor customer first

use result to compute loan (depositor customer)

Strategy 2.

Computer loan depositor first,

and then join the result with customer.

Strategy 3.

Perform the pair of joins at once.

Build an index on loan for loan-number, and on customer for customer-name.

For each tuple t in depositor, look up the corresponding tuples in customer and the corresponding tuples in loan.

Each tuple of deposit is examined exactly once.

Strategy 1 & 2 may use materialization or pipelining

Strategy 3 combines two operations into one special-purpose operation that is more efficient than implementing two joins of two relations.



13.1 Overview



13.4 Sorting

13.5 Join Operation



13.8 Summary


Chapter 13. Summary (1)Chapter 13. Summary (1)

The first action that the system must perform on a query is to translate the query into its internal form, which (for relational database systems) is usually based on the relational algebra.

In the process of generating the internal form of the query, the parser checks the syntax of the user’s query, verifies that the relation names appearing in the query are names of relations in the database, and so on.

If the query was expressed in terms of a view, the parser replaces all references to the view name with the relational-algebra expression to compute the view.

Given a query, there are generally a variety of methods for computing the answer.

It is the responsibility of the query optimizer to transform the query as entered by the user into an equivalent query that can be computed more efficiently.

Chapter 14 covers query optimization.


Chapter 13. Summary (2)Chapter 13. Summary (2) We can process simple selection operations by performing a linear scan, by doing

a binary search, or by making use of indices.

We can handle complex selections by computing unions and intersections of the results of simple selections.

We can sort relations larger than memory by the external merge-sort algorithm.

Queries involving a natural join may be processed in several ways, depending on the availability of indices and the form of physical storage for the relations.

If the join result is almost as large as the Cartesian product of the two relations, a block nested-loop join strategy may be advantageous.

If indices are available, the indexed nested-loop join can be used.

If the relations are sorted, a merge join may be desirable. It may be advantageous to sort a relation prior to join computation (so as to allow use of the merge join strategy).

The hash join algorithm partitions the relations into several pieces, such that each piece of one of the relations fits in memory. The partitioning is carried out with a hash function on the join attributes, so that corresponding pairs of partitions can be joined independently.


Chapter 13. Summary (3)Chapter 13. Summary (3)

Duplicate elimination, projection, set operations (union, intersection and difference), and aggregation can be done by sorting or by hashing.

Outer join operations can be implemented by simple extensions of join algorithms.

Hashing and sorting are dual, in the sense that any operation such as duplicate elimination, projection, aggregation, join, and outer join that can be implemented by hashing can also be implemented by sorting, and vice versa

that is, any operation that can be implemented by sorting can also be implemented by hashing.

An expression can be evaluated by means of materialization, where the system computes the result of each subexpression and stores it on disk, and then uses it to compute the result of the parent expression.

Pipelining helps to avoid writing the results of many subexpressions to disk, by using the results in the parent expression even as they are being generated.


Ch13. Bibliographical Notes (1)Ch13. Bibliographical Notes (1) A query processor must parse statements in the query language, and must

translate them into an internal form. Parsing of query languages differs little from parsing of traditional programming

languages.

Most compiler texts, such as Aho et al. [1986], cover the main parsing techniques,

and present optimization from a programming-language point of view.

Knuth [1973] presents an excellent description of external sorting algorithms, including an optimization that can create initial runs that are (on the average) twice the size of memory. Based on performance studies conducted in the mid-1970s, database systems

of that period used only nested-loop join and merge join. These studies, which were related to the development of System R, determined

that either the nested-loop join or merge join nearly always provided the optimal join method(Blasgen and Eswaran [1976]); hence, these two wer the only join algorithms implemented in System R.

The System R study, however, did not include an analysis of hash join algorithms. Today, hash joins are considered to be highly efficient.


Ch13. Bibliographical Notes (2)Ch13. Bibliographical Notes (2)

Hash join algorithms were initially developed for parallel database systems. Hash join techniques are described in Kitsuregawa et al. [1983], and

extensions including hybrid hash join are described in Shapiro [1986]. Zeller and Gray [1990] and Davison and Graefe [1994] describe hash join

techniques that can adapt to the available memory, which is important in systems where multiple queries may be running at the same time.

Graefe et al. [1998] describes the use of hash joins and hash teams, which allow pipelining of hash-joins by using the same partitioning for all hash-joins in a pipeline sequence, in the Microsoft SQL Server.

Graefe [1993] presents an excellent survey of query-evaluation techniques. An earlier survey of query-processing techniques appears in Jarke and Koch

[1984]. Query processing in main memory database is covered by DeWitt et al. [1984]

and Whang and Krishnamurthy [1990]. Kim [1982] and Kim [1984] describe join strategies and the optimal use of

available main memory.



13.1 Overview



13.4 Sorting

13.5 Join Operation



13.8 Summary

Database System Concepts, 5th Ed.

©Silberschatz, Korth and SudarshanSee www.db-book.com for conditions on re-use

End of ChapterEnd of Chapter

chapter 13: query processing

Documents