VNSISPL_DBMS_Concepts_ch12

Chapter 12: Indexing and Hashing

Database System Concepts, 1 st Ed.
© VNS InfoSolutions Private Limited, Varanasi(UP), India 221002
See www.vnsispl.com for conditions on re-use

Chapter 12: Indexing and Hashing
s Basic Concepts
s Ordered Indices
s B+-Tree Index Files
s B-Tree Index Files
s Static Hashing
s Dynamic Hashing
s Comparison of Ordered Indexing and Hashing
s Index Definition in SQL
s Multiple-Key Access

Database System Concepts – 1 st Ed. 12.2 ©VNS InfoSolutions Private Limited, Varanasi(UP), India 22100

Basic Concepts
s Indexing mechanisms used to speed up access to desired data.
q E.g., author catalog in library
s Search Key - attribute to set of attributes used to look up records in a
file.
s An index file consists of records (called index entries) of the form

search-key pointer
s Index files are typically much smaller than the original file
s Two basic kinds of indices:
q Ordered indices: search keys are stored in sorted order
q Hash indices: search keys are distributed uniformly across
“buckets” using a “hash function”.


Index Evaluation Metrics
s Access types supported efficiently. E.g.,
q records with a specified value in the attribute
q or records with an attribute value falling in a specified range of
values.
s Access time
s Insertion time
s Deletion time
s Space overhead


Ordered Indices

s In an ordered index, index entries are stored sorted on the search
key value. E.g., author catalog in library.
s Primary index: in a sequentially ordered file, the index whose search
key specifies the sequential order of the file.
q Also called clustering index
q The search key of a primary index is usually but not necessarily the
primary key.
s Secondary index: an index whose search key specifies an order
different from the sequential order of the file. Also called
non-clustering index.
s Index-sequential file: ordered sequential file with a primary index.


Dense Index Files
s Dense index — Index record appears for every search-key value in
the file.


Sparse Index Files
s Sparse Index: contains index records for only some search-key
values.
q Applicable when records are sequentially ordered on search-key
s To locate a record with search-key value K we:
q Find index record with largest search-key value < K
q Search file sequentially starting at the record to which the index
record points


Sparse Index Files (Cont.)
s Compared to dense indices:
q Less space and less maintenance overhead for insertions and
deletions.
q Generally slower than dense index for locating records.
s Good tradeoff: sparse index with an index entry for every block in
file, corresponding to least search-key value in the block.


Multilevel Index
s If primary index does not fit in memory, access becomes
expensive.
s Solution: treat primary index kept on disk as a sequential file
and construct a sparse index on it.
q outer index – a sparse index of primary index
q inner index – the primary index file
s If even outer index is too large to fit in main memory, yet
another level of index can be created, and so on.
s Indices at all levels must be updated on insertion or deletion
from the file.


Multilevel Index (Cont.)


Index Update: Deletion
s If deleted record was the only record in the file with its particular search-
key value, the search-key is deleted from the index also.
s Single-level index deletion:
q Dense indices – deletion of search-key:similar to file record
deletion.
q Sparse indices –
 if an entry for the search key exists in the index, it is deleted by
replacing the entry in the index with the next search-key value in
the file (in search-key order).
 If the next search-key value already has an index entry, the entry
is deleted instead of being replaced.


Index Update: Insertion
s Single-level index insertion:
q Perform a lookup using the search-key value appearing in the
record to be inserted.
q Dense indices – if the search-key value does not appear in the
index, insert it.
q Sparse indices – if index stores an entry for each block of the
file, no change needs to be made to the index unless a new block
is created.
 If a new block is created, the first search-key value appearing
in the new block is inserted into the index.
s Multilevel insertion (as well as deletion) algorithms are simple
extensions of the single-level algorithms


Secondary Indices
s Frequently, one wants to find all the records whose values in a
certain field (which is not the search-key of the primary index) satisfy
some condition.
q Example 1: In the account relation stored sequentially by
account number, we may want to find all accounts in a particular
branch
q Example 2: as above, but where we want to find all accounts
with a specified balance or range of balances
s We can have a secondary index with an index record for each
search-key value


Secondary Indices Example

Secondary index on balance field of account

s Index record points to a bucket that contains pointers to all the
actual records with that particular search-key value.
s Secondary indices have to be dense


Primary and Secondary Indices
s Indices offer substantial benefits when searching for records.
s BUT: Updating indices imposes overhead on database modification --
when a file is modified, every index on the file must be updated,
s Sequential scan using primary index is efficient, but a sequential scan
using a secondary index is expensive
q Each record access may fetch a new block from disk
q Block fetch requires about 5 to 10 milliseconds
 versus about 100 nanoseconds for memory access


B+ -Tree Index Files
B+-tree indices are an alternative to indexed-sequential files.

s Disadvantage of indexed-sequential files
q performance degrades as file grows, since many overflow blocks
get created.
q Periodic reorganization of entire file is required.
s Advantage of B+-tree index files:
q automatically reorganizes itself with small, local, changes, in the
face of insertions and deletions.
q Reorganization of entire file is not required to maintain
performance.
s (Minor) disadvantage of B+-trees:
q extra insertion and deletion overhead, space overhead.
s Advantages of B+-trees outweigh disadvantages
q B+-trees are used extensively


B+ -Tree Index Files (Cont.)
A B+-tree is a rooted tree satisfying the following properties:
s All paths from root to leaf are of the same length
s Each node that is not a root or a leaf has between n/2 and n
children.
s A leaf node has between (n–1)/2 and n–1 values
s Special cases:
q If the root is not a leaf, it has at least 2 children.
q If the root is a leaf (that is, there are no other nodes in the
tree), it can have between 0 and (n–1) values.


B+ -Tree Node Structure
s Typical node

q Ki are the search-key values
q Pi are pointers to children (for non-leaf nodes) or pointers to
records or buckets of records (for leaf nodes).
s The search-keys in a node are ordered
K1 < K2 < K3 < . . . < Kn– 1


Leaf Nodes in B+ -Trees
Properties of a leaf node:
s For i = 1, 2, . . ., n– 1, pointer Pi either points to a file record with search-
key value Ki, or to a bucket of pointers to file records, each record
having search-key value Ki. Only need bucket structure if search-key
does not form a primary key.
s If Li, Lj are leaf nodes and i < j, Li’s search-key values are less than Lj’s
search-key values
s Pn points to next leaf node in search-key order


Non-Leaf Nodes in B+ -Trees
s Non leaf nodes form a multi-level sparse index on the leaf nodes. For
a non-leaf node with m pointers:
q All the search-keys in the subtree to which P1 points are less than
K1
q For 2 ≤ i ≤ n – 1, all the search-keys in the subtree to which Pi
points have values greater than or equal to Ki–1 and less than Ki
q All the search-keys in the subtree to which Pn points have values
greater than or equal to Kn–1


Example of a B+ -tree

B+-tree for account file (n = 3)


Example of B+ -tree

B+-tree for account file (n = 5)

s Leaf nodes must have between 2 and 4 values
((n–1)/2 and n –1, with n = 5).
s Non-leaf nodes other than root must have between 3 and 5
children ((n/2 and n with n =5).
s Root must have at least 2 children.


Observations about B+ -trees
s Since the inter-node connections are done by pointers, “logically”
close blocks need not be “physically” close.
s The non-leaf levels of the B+-tree form a hierarchy of sparse indices.
s The B+-tree contains a relatively small number of levels
 Level below root has at least 2* n/2 values
 Next level has at least 2* n/2 * n/2 values
 .. etc.
q If there are K search-key values in the file, the tree height is no
more than  logn/2(K)
q thus searches can be conducted efficiently.
s Insertions and deletions to the main file can be handled efficiently, as
the index can be restructured in logarithmic time (as we shall see).


Queries on B+ -Trees
s Find all records with a search-key value of k.
q N=root
q Repeat
Examine N for the smallest search-key value > k.
If such a value exists, assume it is Ki. Then set N = Pi
Otherwise k ≥ Kn–1. Set N = Pn
Until N is a leaf node
q If for some i, key Ki = k follow pointer Pi to the desired record or bucket.
q Else no record with search-key value k exists.


Queries on B+- Trees (Cont.)
s If there are K search-key values in the file, the height of the tree is no
more than logn/2(K).
s A node is generally the same size as a disk block, typically 4
kilobytes
q and n is typically around 100 (40 bytes per index entry).
s With 1 million search key values and n = 100
q at most log50(1,000,000) = 4 nodes are accessed in a lookup.
s Contrast this with a balanced binary tree with 1 million search key
values — around 20 nodes are accessed in a lookup
q above difference is significant since every node access may need
a disk I/O, costing around 20 milliseconds


Updates on B+ -Trees: Insertion

1. Find the leaf node in which the search-key value would appear
2. If the search-key value is already present in the leaf node
q Add record to the file
q If necessary add a pointer to the bucket.
3. If the search-key value is not present, then
1. add the record to the main file (and create a bucket if
necessary)
2. If there is room in the leaf node, insert (key-value, pointer)
pair in the leaf node
3. Otherwise, split the node (along with the new (key-value,
pointer) entry) as discussed in the next slide.


Updates on B +-Trees: Insertion
(Cont.)
s Splitting a leaf node:
q take the n (search-key value, pointer) pairs (including the one
being inserted) in sorted order. Place the first n/2 in the original
node, and the rest in a new node.
q let the new node be p, and let k be the least key value in p. Insert
(k,p) in the parent of the node being split.
q If the parent is full, split it and propagate the split further up.
s Splitting of nodes proceeds upwards till a node that is not full is found.
q In the worst case the root node may be split increasing the height
of the tree by 1.

Result of splitting node containing Brighton and Downtown on inserting Clearview
Next step: insert entry with (Downtown,pointer-to-new-node) into parent


Updates on B -Trees: Insertion
(Cont.)

B+-Tree before and after insertion of “Clearview”


Insertion in B+ -Trees (Cont.)
s Splitting a non-leaf node: when inserting (k,p) into an already full
internal node N
q Copy N to an in-memory area M with space for n+1 pointers and n
keys
q Insert (k,p) into M
q Copy P1,K1, …, K n/2-1,P n/2 from M back into node N
q Copy Pn/2+1,K n/2+1,…,Kn,Pn+1 from M into newly allocated node N’
q Insert (K n/2,N’) into parent N
s Read pseudocode in book!

Mianus

Downtown Mianus Perryridge Downtown Redwood


Updates on B+ -Trees: Deletion
s Find the record to be deleted, and remove it from the main file and
from the bucket (if present)
s Remove (search-key value, pointer) from the leaf node if there is no
bucket or if the bucket has become empty
s If the node has too few entries due to the removal, and the entries in
the node and a sibling fit into a single node, then merge siblings:
q Insert all the search-key values in the two nodes into a single node
(the one on the left), and delete the other node.
q Delete the pair (Ki– 1, Pi), where Pi is the pointer to the deleted
node, from its parent, recursively using the above procedure.


Updates on B+ -Trees: Deletion
s Otherwise, if the node has too few entries due to the removal, but the
entries in the node and a sibling do not fit into a single node, then
redistribute pointers:
q Redistribute the pointers between the node and a sibling such that
both have more than the minimum number of entries.
q Update the corresponding search-key value in the parent of the
node.
s The node deletions may cascade upwards till a node which has n/2
or more pointers is found.
s If the root node has only one pointer after deletion, it is deleted and
the sole child becomes the root.


Examples of B + -Tree Deletion

Before and after deleting “Downtown”
s Deleting “Downtown” causes merging of under-full leaves
q leaf node can become empty only for n=3!


Examples of B +-Tree Deletion
(Cont.)

Deletion of “Perryridge” from result of previous example
s Leaf with “Perryridge” becomes underfull (actually empty, in this special case) and
merged with its sibling.
s As a result “Perryridge” node’s parent became underfull, and was merged with its sibling
q Value separating two nodes (at parent) moves into merged node
q Entry deleted from parent
s Root node then has only one child, and is deleted


Example of B+ -tree Deletion (Cont.)

Before and after deletion of “Perryridge” from earlier example
s Parent of leaf containing Perryridge became underfull, and borrowed a
pointer from its left sibling
s Search-key value in the parent’s parent changes as a result

B+ -Tree File Organization
s Index file degradation problem is solved by using B+-Tree indices.
s Data file degradation problem is solved by using B+-Tree File
Organization.
s The leaf nodes in a B+-tree file organization store records, instead of
pointers.
s Leaf nodes are still required to be half full
q Since records are larger than pointers, the maximum number of
records that can be stored in a leaf node is less than the number of
pointers in a nonleaf node.
s Insertion and deletion are handled in the same way as insertion and
deletion of entries in a B+-tree index.


B+ -Tree File Organization (Cont.)

Example of B+-tree File Organization
s Good space utilization important since records use more space than
pointers.
s To improve space utilization, involve more sibling nodes in redistribution
during splits and merges
q Involving 2 siblings in redistribution (to avoid split / merge where
possible) results in each node having at least  2n / 3 entries


Indexing Strings
s Variable length strings as keys
q Variable fanout
q Use space utilization as criterion for splitting, not number of
pointers
s Prefix compression
q Key values at internal nodes can be prefixes of full key
 Keep enough characters to distinguish entries in the subtrees
separated by the key value
– E.g. “Silas” and “Silberschatz” can be separated by “Silb”
q Keys in leaf node can be compressed by sharing common prefixes


B-Tree Index Files
s Similar to B+-tree, but B-tree allows search-key values to
appear only once; eliminates redundant storage of search
keys.
s Search keys in nonleaf nodes appear nowhere else in the B-
tree; an additional pointer field for each search key in a
nonleaf node must be included.
s Generalized B-tree leaf node

s Nonleaf node – pointers Bi are the bucket or file record
pointers.


B-Tree Index File Example

B-tree (above) and B+-tree (below) on same data


B-Tree Index Files (Cont.)
s Advantages of B-Tree indices:
q May use less tree nodes than a corresponding B+-Tree.
q Sometimes possible to find search-key value before reaching leaf
node.
s Disadvantages of B-Tree indices:
q Only small fraction of all search-key values are found early
q Non-leaf nodes are larger, so fan-out is reduced. Thus, B-Trees
typically have greater depth than corresponding B+-Tree
q Insertion and deletion more complicated than in B+-Trees
q Implementation is harder than B+-Trees.
s Typically, advantages of B-Trees do not out weigh disadvantages.


Multiple-Key Access
s Use multiple indices for certain types of queries.
s Example:
select account_number
from account
where branch_name = “Perryridge” and balance = 1000
s Possible strategies for processing query using indices on single
attributes:
1. Use index on branch_name to find accounts with branch name
Perryridge; test balance = 1000
2. Use index on balance to find accounts with balances of $1000;
test branch_name = “Perryridge”.
3. Use branch_name index to find pointers to all records pertaining to
the Perryridge branch. Similarly use index on balance. Take
intersection of both sets of pointers obtained.


Indices on Multiple Keys
s Composite search keys are search keys containing more than one
attribute
q E.g. (branch_name, balance)
s Lexicographic ordering: (a1, a2) < (b1, b2) if either
q a1 < b1, or
q a1=b1 and a2 < b2


Indices on Multiple Attributes

Suppose we have an index on combined search-key
(branch_name, balance).

s With the where clause
where branch_name = “Perryridge” and balance = 1000
the index on (branch_name, balance) can be used to fetch only
records that satisfy both conditions.
q Using separate indices in less efficient — we may fetch many
records (or pointers) that satisfy only one of the conditions.
s Can also efficiently handle
where branch_name = “Perryridge” and balance < 1000
s But cannot efficiently handle
where branch_name < “Perryridge” and balance = 1000
q May fetch many records that satisfy the first but not the second
condition


Non-Unique Search Keys
s Alternatives:
q Buckets on separate block (bad idea)
q List of tuple pointers with each key
 Extra code to handle long lists
 Deletion of a tuple can be expensive if there are many
duplicates on search key (why?)
 Low space overhead, no extra cost for queries
q Make search key unique by adding a record-identifier
 Extra storage overhead for keys
 Simpler code for insertion/deletion
 Widely used


Other Issues in Indexing
s Covering indices
q Add extra attributes to index so (some) queries can avoid fetching
the actual records
 Particularly useful for secondary indices
– Why?
q Can store extra attributes only at leaf
s Record relocation and secondary indices
q If a record moves, all secondary indices that store record pointers
have to be updated
q Node splits in B+-tree file organizations become very expensive
q Solution: use primary-index search key instead of record pointer in
secondary index
 Extra traversal of primary index to locate record
– Higher cost for queries, but node splits are cheap
 Add record-id if primary-index search key is non-unique


Hashing


Static Hashing

s A bucket is a unit of storage containing one or more records (a
bucket is typically a disk block).
s In a hash file organization we obtain the bucket of a record
directly from its search-key value using a hash function.
s Hash function h is a function from the set of all search-key values K
to the set of all bucket addresses B.
s Hash function is used to locate records for access, insertion as well
as deletion.
s Records with different search-key values may be mapped to the
same bucket; thus entire bucket has to be searched sequentially to
locate a record.


Example of Hash File Organization

Hash file organization of account file, using branch_name as key
(See figure in next slide.)

s There are 10 buckets,
s The binary representation of the ith character is assumed to be the
integer i.
s The hash function returns the sum of the binary representations of
the characters modulo 10
q E.g. h(Perryridge) = 5 h(Round Hill) = 3 h(Brighton) = 3


Example of Hash File Organization

Hash file organization
of account file, using
branch_name as key
(see previous slide for
details).


Hash Functions
s Worst hash function maps all search-key values to the same bucket;
this makes access time proportional to the number of search-key values
in the file.
s An ideal hash function is uniform, i.e., each bucket is assigned the
same number of search-key values from the set of all possible values.
s Ideal hash function is random, so each bucket will have the same
number of records assigned to it irrespective of the actual distribution of
search-key values in the file.
s Typical hash functions perform computation on the internal binary
representation of the search-key.
q For example, for a string search-key, the binary representations of
all the characters in the string could be added and the sum modulo
the number of buckets could be returned. .


Handling of Bucket Overflows
s Bucket overflow can occur because of
q Insufficient buckets
q Skew in distribution of records. This can occur due to two
reasons:
 multiple records have same search-key value
 chosen hash function produces non-uniform distribution of key
values
s Although the probability of bucket overflow can be reduced, it cannot
be eliminated; it is handled by using overflow buckets.


Handling of Bucket Overflows
(Cont.)

s Overflow chaining – the overflow buckets of a given bucket are chained
together in a linked list.
s Above scheme is called closed hashing.
q An alternative, called open hashing, which does not use overflow
buckets, is not suitable for database applications.


Hash Indices
s Hashing can be used not only for file organization, but also for index-
structure creation.
s A hash index organizes the search keys, with their associated record
pointers, into a hash file structure.
s Strictly speaking, hash indices are always secondary indices
q if the file itself is organized using hashing, a separate primary hash
index on it using the same search-key is unnecessary.
q However, we use the term hash index to refer to both secondary
index structures and hash organized files.


Example of Hash Index


Deficiencies of Static Hashing
s In static hashing, function h maps search-key values to a fixed set of B
of bucket addresses. Databases grow or shrink with time.
q If initial number of buckets is too small, and file grows, performance
will degrade due to too much overflows.
q If space is allocated for anticipated growth, a significant amount of
space will be wasted initially (and buckets will be underfull).
q If database shrinks, again space will be wasted.
s One solution: periodic re-organization of the file with a new hash
function
q Expensive, disrupts normal operations
s Better solution: allow the number of buckets to be modified dynamically.


Dynamic Hashing
s Good for database that grows and shrinks in size
s Allows the hash function to be modified dynamically
s Extendable hashing – one form of dynamic hashing
q Hash function generates values over a large range — typically b-bit
integers, with b = 32.
q At any time use only a prefix of the hash function to index into a
table of bucket addresses.
q Let the length of the prefix be i bits, 0 ≤ i ≤ 32.
 Bucket address table size = 2i. Initially i = 0
 Value of i grows and shrinks as the size of the database grows
and shrinks.
q Multiple entries in the bucket address table may point to a bucket
(why?)
q Thus, actual number of buckets is < 2i
 The number of buckets also changes dynamically due to
coalescing and splitting of buckets.


General Extendable Hash Structure

In this structure, i2 = i3 = i, whereas i1 = i – 1 (see next
slide for details)


Use of Extendable Hash Structure
s Each bucket j stores a value ij
q All the entries that point to the same bucket have the same values on
the first ij bits.
s To locate the bucket containing search-key Kj:

1. Compute h(Kj) = X
2. Use the first i high order bits of X as a displacement into bucket
address table, and follow the pointer to appropriate bucket
s To insert a record with search-key value Kj
q follow same procedure as look-up and locate the bucket, say j.
q If there is room in the bucket j insert record in the bucket.
q Else the bucket must be split and insertion re-attempted (next slide.)
 Overflow buckets used instead in some cases (will see shortly)


Insertion in Extendable Hash Structure
(Cont)
To split a bucket j when inserting record with search-key value Kj:

s If i > ij (more than one pointer to bucket j)
q allocate a new bucket z, and set ij = iz = (ij + 1)
q Update the second half of the bucket address table entries originally
pointing to j, to point to z
q remove each record in bucket j and reinsert (in j or z)
q recompute new bucket for Kj and insert record in the bucket (further
splitting is required if the bucket is still full)
s If i = ij (only one pointer to bucket j)
q If i reaches some limit b, or too many splits have happened in this
insertion, create an overflow bucket
q Else
 increment i and double the size of the bucket address table.
 replace each entry in the table by two entries that point to the
same bucket.
 recompute new bucket address table entry for Kj
Now i > ij so use the first case above.

Deletion in Extendable Hash
Structure
s To delete a key value,
q locate it in its bucket and remove it.
q The bucket itself can be removed if it becomes empty (with
appropriate updates to the bucket address table).
q Coalescing of buckets can be done (can coalesce only with a
“buddy” bucket having same value of ij and same ij –1 prefix, if it is
present)
q Decreasing bucket address table size is also possible
 Note: decreasing bucket address table size is an expensive
operation and should be done only if number of buckets becomes
much smaller than the size of the table


Use of Extendable Hash Structure:
Example

Initial Hash structure, bucket size = 2


Example (Cont.)
s Hash structure after insertion of one Brighton and two Downtown
records


Example (Cont.)
Hash structure after insertion of Mianus record


Example (Cont.)

Hash structure after insertion of three Perryridge records


Example (Cont.)
s Hash structure after insertion of Redwood and Round Hill records


Extendable Hashing vs. Other
Schemes
s Benefits of extendable hashing:
q Hash performance does not degrade with growth of file
q Minimal space overhead
s Disadvantages of extendable hashing
q Extra level of indirection to find desired record
q Bucket address table may itself become very big (larger than
memory)
 Cannot allocate very large contiguous areas on disk either
 Solution: B+-tree structure to locate desired record in bucket
address table
q Changing size of bucket address table is an expensive operation
s Linear hashing is an alternative mechanism
q Allows incremental growth of its directory (equivalent to bucket
address table)
q At the cost of more bucket overflows


Comparison of Ordered Indexing and Hashing

s Cost of periodic re-organization
s Relative frequency of insertions and deletions
s Is it desirable to optimize average access time at the expense of
worst-case access time?
s Expected type of queries:
q Hashing is generally better at retrieving records having a specified
value of the key.
q If range queries are common, ordered indices are to be preferred
s In practice:
q PostgreSQL supports hash indices, but discourages use due to
poor performance
q Oracle supports static hash organization, but not hash indices
q SQLServer supports only B+-trees


Bitmap Indices
s Bitmap indices are a special type of index designed for efficient
querying on multiple keys
s Records in a relation are assumed to be numbered sequentially from,
say, 0
q Given a number n it must be easy to retrieve record n
 Particularly easy if records are of fixed size
s Applicable on attributes that take on a relatively small number of
distinct values
q E.g. gender, country, state, …
q E.g. income-level (income broken up into a small number of levels
such as 0-9999, 10000-19999, 20000-50000, 50000- infinity)
s A bitmap is simply an array of bits


Bitmap Indices (Cont.)
s In its simplest form a bitmap index on an attribute has a bitmap for
each value of the attribute
q Bitmap has as many bits as records
q In a bitmap for value v, the bit for a record is 1 if the record has the
value v for the attribute, and is 0 otherwise


s Bitmap indices are useful for queries on multiple attributes
q not particularly useful for single attribute queries
s Queries are answered using bitmap operations
q Intersection (and)
q Union (or)
q Complementation (not)
s Each operation takes two bitmaps of the same size and applies the
operation on corresponding bits to get the result bitmap
q E.g. 100110 AND 110011 = 100010
100110 OR 110011 = 110111
NOT 100110 = 011001
q Males with income level L1: 10010 AND 10100 = 10000
 Can then retrieve required tuples.
 Counting number of matching tuples is even faster


s Bitmap indices generally very small compared with relation size
q E.g. if record is 100 bytes, space for a single bitmap is 1/800 of space
used by relation.
 If number of distinct attribute values is 8, bitmap is only 1% of
relation size
s Deletion needs to be handled properly
q Existence bitmap to note if there is a valid record at a record location
q Needed for complementation
 not(A=v): (NOT bitmap-A-v) AND ExistenceBitmap
s Should keep bitmaps for all values, even null value
q To correctly handle SQL null semantics for NOT(A=v):
 intersect above result with (NOT bitmap-A-Null)


Efficient Implementation of Bitmap
Operations

s Bitmaps are packed into words; a single word and (a basic CPU
instruction) computes and of 32 or 64 bits at once
q E.g. 1-million-bit maps can be and-ed with just 31,250 instruction
s Counting number of 1s can be done fast by a trick:
q Use each byte to index into a precomputed array of 256 elements
each storing the count of 1s in the binary representation
 Can use pairs of bytes to speed up further at a higher memory
cost
q Add up the retrieved counts
s Bitmaps can be used instead of Tuple-ID lists at leaf levels of
B+-trees, for values that have a large number of matching records
q Worthwhile if > 1/64 of the records have that value, assuming a
tuple-id is 64 bits
q Above technique merges benefits of bitmap and B+-tree indices


Index Definition in SQL
s Create an index
create index <index-name> on <relation-name>
(<attribute-list>)
E.g.: create index b-index on branch(branch_name)
s Use create unique index to indirectly specify and enforce the
condition that the search key is a candidate key is a candidate key.
q Not really required if SQL unique integrity constraint is supported
s To drop an index
drop index <index-name>
s Most database systems allow specification of type of index, and
clustering.


End of Chapter


Partitioned Hashing
s Hash values are split into segments that depend on each
attribute of the search-key.
(A1, A2, . . . , An) for n attribute search-key
s Example: n = 2, for customer, search-key being
(customer-street, customer-city)
search-key value hash value
(Main, Harrison) 101 111
(Main, Brooklyn) 101 001
(Park, Palo Alto) 010 010
(Spring, Brooklyn) 001 001
(Alma, Palo Alto) 110 010
s To answer equality query on single attribute, need to look up
multiple buckets. Similar in effect to grid files.


Sequential File For account Records


Sample account File


Figure 12.2


Figure 12.14


Figure 12.25


Grid Files
s Structure used to speed the processing of general multiple search-
key queries involving one or more comparison operators.
s The grid file has a single grid array and one linear scale for each
search-key attribute. The grid array has number of dimensions
equal to number of search-key attributes.
s Multiple cells of grid array can point to same bucket
s To find the bucket for a search-key value, locate the row and column
of its cell using the linear scales and follow pointer


Example Grid File for account


Queries on a Grid File
s A grid file on two attributes A and B can handle queries of all following
forms with reasonable efficiency
q (a1 ≤ A ≤ a2)
q (b1 ≤ B ≤ b2)
q (a1 ≤ A ≤ a2 ∧ b1 ≤ B ≤ b2),.
s E.g., to answer (a1 ≤ A ≤ a2 ∧ b1 ≤ B ≤ b2), use linear scales to find
corresponding candidate grid array cells, and look up all the buckets
pointed to from those cells.


Grid Files (Cont.)
s During insertion, if a bucket becomes full, new bucket can be created
if more than one cell points to it.
q Idea similar to extendable hashing, but on multiple dimensions
q If only one cell points to it, either an overflow bucket must be
created or the grid size must be increased
s Linear scales must be chosen to uniformly distribute records across
cells.
q Otherwise there will be too many overflow buckets.
s Periodic re-organization to increase grid size will help.
q But reorganization can be very expensive.
s Space overhead of grid array can be high.
s R-trees (Chapter 23) are an alternative


VNSISPL_DBMS_Concepts_ch12

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