algorithms in bioinformatics: a practical introduction

Algorithms in Bioinformatics: A Practical Introduction

Suffix tree

Overview What is suffix tree? Simple application of suffix tree Linear time algorithm for

constructing suffix tree Suffix array FM-index 1-mismatch search

Suffix Trie

E.g. consider the string S = acacag$Suffix Trie: a ties of all possible suffices of S

$ a c

a

g

c

a

g

c g

7

c g

1

3

5

2

4

6

a

gg

a

$7g$6ag$5cag$4acag$3cacag$2acacag$1Suffix

$

$

$ $

$

$

Suffix TreeSuffix tree for S=acacag$: merge nodes with only one

child 1 2 3 4 5 6 7a c a c a g $S=

v

$ a ca

g

ca

g

cg

c g

7

1 3

5

2 4

6

ag

g$ $

$

$

$$a

Path-label of node v is “aca”Denoted as (v)

“ca” is an edge label

This is a leaf edge

Size of Suffix Tree (I) How big is a suffix tree?

Suffix tree has exactly n leaves and at most 2n-1 edges The total length of all edge labels is O(n2).

Can we store suffix tree using o(n2) bit space?1 2 3 4 5 6 7a c a c a g $S=

$ a ca

g

ca

gc

gc

g

7

1 3

5

2 4

6

ag

g$ $

$

$

$$a

Size of Suffix Tree (II)Suffix tree has exactly n leaves and at most 2n-1 edgesNote that each edge label can be represented using 2

indicesThus, suffix tree can be represented using O(n log n) bits

1 2 3 4 5 6 7a c a c a g $S=$ a c

ag

ca

gc

gc

g

7

1 3

5

2 4

6

ag

g$ $

$

$

$$

6,7

6,76,7

6,74,7

4,7

a

7,71,1

2,32,3

Note: The end index of every leaf edge should be 7, the last index of S. Thus, for leaf edges, we only need to store the start index.

Property of suffix tree Fact: For any internal node v in the

suffix tree, if the path label of v is (v)=ap, then there exists another node w in the suffix

tree such that (w)=p.

Proof: Skip the proof.

Definition of Suffix Link: For any internal node v, define its suffix link

sl(v) = w.

Suffix Link example S=acacag$

$ a ca

g

ca

g

cg

c g

7

1 3

5

2 4

6

ag

g$ $

$

$

$$a

Generalized suffix tree Build a suffix tree for two or more strings E.g. S1 = acgat#, S2 = cgt$

a c gg t

at t a t

c

41 2 1 3 2 3

tat $ $$

#

g t#

5

##

46

# $

#

Applications of Suffix Tree

Exact string matching problem To find all occurrences of Q in S (searching)

Search for the node x in the suffix tree which represent Q All the leaves in the subtree rooted at x are the

occurrences Time: O(|Q| + occ) where occ is the total no. of

occurrences$ a c

ag

cag

cg

c g

7

1 3

5

2 4

6

ag

g$ $

$

$

$$a

E.g. S = acacag$Q = acaOccurrences: 1, 3

Longest repeated substring problem To find the longest repeated substring in S

Find the deepest internal node Time: O(n)

$ a ca

g

cag

cg

c g

7

1 3

5

2 4

6

ag

g$ $

$

$

$$a

E.g. S = acacag$The longest repeat is aca.

Longest common substring problem To find the longest common substring of

two or more sequences Note: 1970, Don Knuth conjectured that a

linear time algorithm for this problem is impossible

Now, we know that it can be solved in linear time.

E.g. consider two string S1 and S2,1. Build generalized suffix tree for S1# and S2$2. Then, mark each internal node with leaves

representing suffixes of both S1 and S2.3. Report the deepest marked node

Example for the longest common substring E.g. S1 = acgat#, S2 = cgt$ The longest common substring is “cg”. Its

length is 2.

a c gg t

at t a t

c

41 2 1 3 2 3

ta $ $$#

g t#

5

##

46

# $

#t

Longest common prefix (I) Given a string S. For any i, j,

Denote lcp(i, j) be the length of the longest common prefix of suffix i and j of S.

$ a ca

g

cag

cg

c g

7

1 3

5

24

6

ag

g$ $

$

$

$$a

1 2 3 4 5 6 7a c a c a g $S=

The longest common prefix of suffix 1 and suffix 3 is aca!

lcp(1, 3) = 3

Longest common prefix (II) Note that the lowest common ancestor(lca) of

leaves i and j identifies the longest common prefix. lcp(i, j) = |(lca(i, j))|. A well-know result:

Consider a tree of size n, after an O(n) time preprocessing, the lca for any two nodes can be returned in O(1) time.

First obtained by Harel and Tarjan (SIAM J. Comp. 1984) Simplified by Schieber and Vishkin (SIAM J. Comp. 1988)

Based on the above result, After an O(n) time preprocessing,

For any suffix i and suffix j, we can compute the longest common prefix of them in O(1) time.

Finding Palindrome (I) Given a string S, palindrome is a substring u of S s.t. u =

ur

E.g. ACAGACA Consider a palindrome u=S[i..i+|u|-1], u is called a

maximal palindrome if S[i’..j’] is not a palindrome for any [i’..j’][i..i+|u|-1].

Note that every palindrome is contained in a maximal palindrome.

Thus, maximal palindromes are a compact way to represent all palindromes.

Complemented Palindrome is a string u s.t. u = ūr

E.g. ACAUGU Maximal complemented palindrome is defined similarly.

Finding Palindrome (II) Recall that restriction enzyme usually is in

the form of complemented palindrome. This motivates the following two problems:

The palindrome problem: Given a string S (representing the genome) of length n,

the problem is to locate all maximal palindromes in S. The complemented palindrome problem:

Given a string S (representing the genome) of length n, the problem is to locate all maximal complemented palindromes in S.

Properties of palindrome (I) If S[i..i+k-1]=Sr[n-i+1..n-i+k], then

u=S[i-k+1..i+k-1] is an odd length palindrome　

1 i n

S a a a b a a a

1 n

Sr a a a b a a a

i-k+1 i+k-1

n-i+1

Properties of palindrome (II) If S[i..i+k-1]=Sr[n-i+2..n-i+k+1],

then u=S[i-k..i+k-1] is an even length palindrome　

1 i-k i n

S a a a b b a a a

1 n-i+2 n

Sr a a a b b a a a

i+k-1

Solution to the palindrome problem Preprocess S and Sr so that any longest

common prefix query can be answered in constant time.

For i=1 to n, Find the longest common prefix for (Si, Sr

n-i+1). If the longest prefix is k, we find an odd length maximal palindrome S[i-k+1..i+k-1].

Find the longest common prefix for (Si, Srn-i+2).

If the longest prefix is k, we find an even length maximal palindrome S[i-k..i+k-1].

Extracting embedded suffix tree from a generalized suffix tree

Input: The generalized suffix tree T of K strings S1, …, SK.

Aim: Compute the suffix tree Ti of the string Si.r

x y

a c gg t

at t a t

c

41 2 1 3 2 3

tat $ $$

#

gzw

t#

5

##

46

# $

#

r

a cgg

tat

ac

41 2 3

tat #

g

w

t#

5

# #

6

#

#

S1 = acgat#, S2 = cgt$T

T1

Extracting embedded suffix tree from a generalized suffix tree Observation: Ti is a subtree of T such that

The leaves of Ti are the leaves of T corresponding to Si. The internal nodes of Ti are the lowest common ancestors

of some leaves for Si. The edges of Ti can be inferred from the ancestor

descendent relationship among those nodes.r

x y

a c gg t

at t a t

c

41 2 1 3 2 3

tat $ $$

#

gzw

t#

5

##

46

# $

#

r

a cgg

tat

ac

41 2 3

tat #

g

w

t#

5

# #

6

#

#

S1 = acgat#, S2 = cgt$T

T1

Extracting embedded suffix tree from a generalized suffix tree

r

a cgg

tat

ac

41 2 3

tat #

g

w

t#

5

# #

6

#

#

r

a cgg

at

ac

41 2 3

tat #

g

w

t# #

6

#

#

r

a cgat

c

41 2

at

g

w

t# #

6

#

#

r

a

c

41

at

g

w

t#

6

#

#

r

ac

1

at

g6

#

#

r

6

#

Common substrings of more than 2 strings (I) Given a set of strings (protein or DNA

sequences), we want to know what substrings are common to a large number of these strings?

Why this question is important? DNA and protein sequences will evolve. If a

substring occur commonly in wide range of species. This may mean that the substring is critical for the correct functionality.

Common substrings of more than 2 strings (II) Given K strings whose total length

is n. For every 2kK, define l(k) be the

length of the longest substring common to at least k of these strings.

The problem is to compute l(k) for all k.

Common substrings of more than 2 strings (III) Example:

Consider a set of 5 strings { sandollar, sandlot, handler, grand, pantry }

Then, we havek l(k

)corresponding substring

2 4 sand

3 3 and

4 3 and

5 2 an

Common substrings of more than 2 strings (IV) Illustrating the solution by example:

S1 = aacg$, S2 = acgc#, S3 = cga%. (K=3)

1. Build a generalized suffix tree T for the K strings in O(n) time.

$ a cg

cgg

c

5

2 1

$#

54

#%

3

%

1

g$

ca

4

# c

3 1

$ #

2

a%

c

4 2

$ #

3

a%

Common substrings of more than 2 strings (V)

2. By traversing T, for each internal node v, compute its string depth. In total, O(n) time.

0

1 1

3

$ a cg

cgg

c

5

2 1

$#

54

#%

3

%

1

g$

ca

4

#2

c

3 1

$ #

2

a%

1

c

4 2

$ #

3

a%

Common substrings of more than 2 strings (VI)3. By traversing T, for each internal node v, compute

C(v). [C(v) is defined as the number of distinct termination symbols in the subtree rooted at v]

This step takes O(Kn) time.3

3 3

2

$ a cg

cgg

c

5

2 1

$#

54

#%

3

%

1

g$

ca

4

#3

c

3 1

$ #

2

a%

3

c

4 2

$ #

3

a%

Common substrings of more than 2 strings (VII)4. Traverse T and visit every internal node v. For each v, if

V(C(v)) < string-depth of v, set V(C(v)) = string-depth of v. [After step 4, V(k) = the length of the longest substring common to exactly k of these strings.]

5. l(k)=V(k). For i=k-1 downto 2, l(i)=max{l(i+1), V(i)}.

This two steps take O(n) time.

For our example, V(2) = 3, V(3) = 2. Thus, l(3) = 2, l(2) = 3.

In total, this algorithm takes O(Kn) time. Actually, we can improve this algorithm to O(n) time by

mean of lcp!

Linear time algorithm for constructing suffix tree

Straightforward construction of suffix tree Consider S = s1s2…sn where sn=$

Algorithm: Initialize the tree with only a root For i = n to 1

Includes S[i..n] into the tree

Time: O(n2)

Example of construction S=acca$

Init For-loop

$

5

I5

a

4

I4

5

$ $a

4

I3

5

$ $

3

a$

c

I2

c

ca a

3 2

a

45

$ $$ $

I1

c

ca a

3 25

$

$ $

cc

a

a

4 1

$$

Construction of generalized suffix tree S’= c#

I1

c

ca a

3 25

$

$ $

cc

a

a

4 1

$$

Init For-loop

J2

c

ca a

3 25

$

$ $

cc

a

a

4 1

$$

2

#

J1

c

ca a

3 25

$

$ $

cc

a

a

4 1

$$

2

##

1

Can we construct a suffix tree in o(n2) time? Yes. We can construct it in O(n) time.

Weiner’s algorithm [1973] Linear time for constant size alphabet, but much space

McGreight’s algorithm [JACM 1976] Linear time for constant size alphabet, quadratic space

Ukkonen’s algorithm [Algorithmica, 1995] Online algorithm, linear time for constant size alphabet, less space

Farach’s algorithm [FOCS 1997] Linear time for general alphabet

Hon,Sadakane, and Sung’s algorithm [FOCS 2003] O(n) bit space O(n logen) time for 0<e<1 O(n) bit space O(n) time for suffix array construction

We will discuss Farach’s algorithm later.

Idea Build Odd Suffix Tree and Even Suffix

Tree Then, merge odd and even suffix tree.

$ a ca

g

cag

cg

c g

7

1 3

5

24

6

ag

g$ $

$

$

$$a

$ a ca

g

cag

cg

c g

7

1 3

5

24

6

ag

g$ $

$

$

$$a

Even Suffix Tree Odd Suffix Tree

Idea Input: a string S of length n1. Recursively compute the suffix tree To of

all suffixes beginning at the odd positions.

To is of size n/2.

2. From To, compute Te which is the suffix tree for all suffixes beginning at the even positions.

3. Merge To and Te to form the suffix tree for S.

Stage 1: Constructing odd suffix tree Given a string S[1..n], we generate a new

string S’[1..n/2] as follows. we map pairs of characters into single

characters as follows: S[1..2], S[3..4], S[5..6], …, S[n-1..n].

Remove the duplicates from the pairs of characters and sort them by radix sort.

S’[i] = rank of S[2i-1..2i] in the sorted list, for i=1, 2, …, n/2.

By recursion, we get the suffix tree T’ for S’

Convert T’ to the odd suffix tree To.

Example (I) S = aaabbbabbaba$ S[1..2]=aa, S[3..4]=ab,

S[5..6]=bb, S[7..8]=ab, S[9..10]=ba, S[11..12]=ba.

By stable sort, aa < ab < ba < bb. Rank(aa)=1, Rank(ab)=2,

Rank(ba)=3, Rank(bb)=4. So, S’=124233$.

Example (II) By recursion, construct the suffix

tree T’ for S’:

7

$

1

12

42

33

$

4

33

$

2

2

24233$

3

6 5

3$$

3

4233$

Example (III) Convert T’ to the odd tree:

13

$

1

aa

ab

bb

ab

ba

ba

$

7

ba

ba

$

ab

3

ba

11 9

ba$

$

5

bbabbaba$

bbbabbaba$

This is not a suffix tree

i 2i-1

Example (IV) Refine the odd tree To:

13

$

1

aa

bb

ba

bb

ab

a$

7

ab

a$

bb

3

11 9

ba$

$

5

babbaba$

bbabbaba$

b

a

a

Time complexity for building the odd tree Let Time(n) be the time to build a

suffix tree for a string of length n.

Stable sorting and refinement of the odd trees take O(n) time.

Build suffix tree for S’ takes Time(n/2). So, Stage 1 takes Time(n/2)+O(n)

time.

Stage 2: Build the even tree1. Generate the lex-ordering of the

leaves in Te.

2. For any two adjacent leaves 2i and 2j, we find lcp(2i, 2j).

3. Construct the even tree Te from left to right (according to the lex-ordering).

Build the even tree (Step 1) We get the lex-ordering of the leaves

in To. Generate the lex-ordering of the

leaves in Te. For each leaf i in To, get the preceding

character c=S[i-1] and form a pair (c,i). Each pair represents a even suffix i-1.

Perform stable sorting on those pairs. We get the lex-ordering of the leaves in Te.

Example Lex-ordering of the leaves in To:

13 < 1 < 7 < 3 < 11 < 9 < 5 The pairs are:

(a,13), ($,1), (b,7), (a, 3), (a, 11), (b, 9), (b, 5). After stable sorting, we have

($, 1), (a, 13), (a, 3), (a, 11), (b, 7), (b, 9), (b, 5).

Hence, the lex-ordering of the leaves of Te: 12 < 2 < 10 < 6 < 8 < 4

S = aaabbbabbaba$

Build the even tree (Step 2) For any two adjacent leaves 2i and 2j,

we first find lcp(2i, 2j).

Observation: lcp(2i, 2j) = lcp(2i+1, 2j+1)+1 if S[2i]=S[2j] 0 otherwise

Proof: If S[2i]S[2j], lcp(2i,2j)=0. Otherwise, lcp(2i,2j)=1+lcp(2i+1,2j+1).

Example Recall that the lex-

ordering of leaves: 12 < 2 < 10 < 6 < 8 < 4.

By the previous observation, we have lcp(8,4)=lcp(9,5)+1=2

Similarly, we have lcp(12,2)=1, lcp(2,10)=1,

lcp(10,6)=0, lcp(6,8)=1, lcp(8,4)=2

13

$

1

aa

bb

ba

bb

ab

a$

7

ab

a$

bb

3

11 9

ba$

$

5

babbaba$

bbabbaba$

b

a

a

Build the even tree (Step 3) Construct the even tree Te from

left to right.

12

a

12

a$

$

2

abbbabbaba$

a

12

$

2

abbbabbaba$

10

ba$

Build the even tree (Step 3)

a

12

$

2

abbbabbaba$

10

ba$

6

babbaba$

a

12

$

2

abbbabbaba$

10

ba$

6

abbaba$

b

8

baba$

a

12

$

2

abbbabbaba$

10

ba$

6

abbaba$

b

8

aba$

4

babbaba$

b

Time complexity for building the even tree

Step 1: O(n) time Step 2: O(n) time Step 3: O(n) time

Stage 3: Merge odd and even trees We can merge To and Te by DFS. However, it takes O(n2) time.

13

$

1

aa

bb

ba

bb

ab

a$

7

ab

a$

bb

3

11 9

ba$

$

5

babbaba$

bbabbaba$

b

a

a

13

$

1

aa

bb

ba

bb

ab

a$

7

ab

a$

bb

3

11 9

ba$

$

5

babbaba$

bbabbaba$

b

a

a

Even tree

Odd tree

a

12

$

2

abbbabbaba$

10

ba$

6

abbaba$

b

8

aba$

4

babbaba$

b

a

12

$

2

abbbabbaba$

10

ba$

6

abbaba$

b

8

aba$

4

babbaba$

b

13$,1

1

7 3

11 9

$,1

b,1

a,1

a,1

12$,1

a,1

2

b,1

10a,2

a,4 b,9

6

b,1

b,5

8

4

b,1

a,2

5b,5

b,8a,11

b,10

a,2a,2

b,1

13$,1

1

7 3

11 9

$,1

b,1

a,1

a,1

12$,1

a,1

2

b,1

10a,2

a,4 b,9

6

b,1

b,5

8

4

b,1

a,2

5b,5

b,8a,2

a,2b,1

Stage 3: Merge odd and even trees We merge To and Te by DFS. We merge two edges as long as they start

with the same character. The merge is ended when one edge is longer than the other.

13

$

1

aa

bb

ba

bb

ab

a$

7

ab

a$

bb

3

11 9

ba$

$

5

babbaba$

bbabbaba$

b

a

aa

12

$

2

abbbabbaba$

10

ba$

6

abbaba$

b

8

aba$

4

babbaba$

b

Merge odd and even trees We merge To and Te by DFS. We merge two edges as long as they start

with the same character. The merge is ended when one edge is longer than the other.

13

$,1

27

3

11 9

$,1

b,1

a,1

a,1

12

$,1

a,11

1

$,1

b,2

10

$,1

a,4b,9

6

b,3

b,3

8

4

b,1

a,4

5

b,3

b,8

Merge odd and even trees The merging may over-merged some nodes. To correct the tree, we need to unmerge some nodes.

13

$,1

27

3

11 9

$,1

b,1

a,1

a,1

12

$,1

a,11

1

$,1

b,2

10

$,1

a,4b,9

6

b,3

b,3

8

4

b,1

a,4

5

b,3

b,8

Definition of L() and d() For every node u which may be over-

merged, there exist two leaves 2i and 2j-1 such that u=lca(2i, 2j-1). Denote L(u) be the correct depth of u, that is,

lcp(2i,2j-1).

Note that lcp(2i,2j-1) = 1+lcp(2i+1,2j) if S[2i]=S[2j-1]; 0 otherwise.

Let v be lca(2i+1,2j). Denote d(u) = v. Note that d() is equivalent to suffix link!

Example of d()

13

$,1

27

3

11 9

$,1

b,1

a,1

a,1

12

$,1

a,11

1$,1

b,2

10

$,1

a,4b,9

6

b,3

b,3

8

4

b,1

a,4

5

b,3

b,8

d()

Relationship between L() and d() Suppose u = lca(2i, 2j-1).

Note1: if u is not the root, then S[2i]=S[2j-1]. Note2: lcp(2i,2j-1) = 1+lcp(2i+1,2j) if S[2i]=S[2j-1]; 0

otherwise Note3: d(u) = lcp(2i+1,2j)

Hence, L(u) = 1 + L(d(u)) if u is not the root. Otherwise, L(u)=0.

Lemma: L(u) = the length of the purple path from u to the root.

Example of L()

13

$,1

27

3

11 9

$,1

b,1

a,1

a,1

12

$,1

a,11

1

$,1

b,2

10

$,1

a,4b,9

6

b,3

b,3

8

4

b,1

a,4

5

b,3

b,8

L( )=2 L( )=3

L( )=4

L( )=2L( )=2

L( )=2

L( )=1L( )=1

Unmerge the border nodes based on L() (I)

13

$,1

27

3

11 9

$,1

b,1

a,1

a,1

12

$,1

a,11

1

$,1

b,2

10

$,1

a,4b,9

6

b,3

b,3

8

4

b,1

a,4

5

b,3

b,8

L( )=2 L( )=3

L( )=4

L( )=2L( )=2

L( )=2

L( )=1L( )=1

Unmerge the border nodes based on L() (II)

13

$,1

1

7 3

11 9

$,1

b,1

a,1

a,1

12

$,1a,1

2

b,1

10

a,2

a,4 b,9

6

b,1

b,5

8

4

b,1

a,2

5

b,5

b,8a,11

b,10

a,2

a,2b,1

Time complexity for merging Merge the tree using DFS takes

O(n) time. Compute the links d() takes O(n)

time. Compute L() takes O(n) time. Unmerge takes O(n) time.

Total time complexity of Farach’s algorithm

Stage 1: Time(n/2)+O(n) Stage 2: O(n) Stage 3: O(n)

Thus, Time(n) = Time(n/2)+O(n). By solving the equation,

Time(n)= O(n).

Disadvantage of suffix tree Suffix tree is space inefficient. It

requires O(n||log n) bits. Manber and Myers (SIAM J. Comp

1993) proposes a new data structure, called suffix array, which has a similar functionality as suffix tree. Moreover, it only requires O(n log n) bits.

Suffix Array (I) It is just sorted suffixes. E.g. consider S = acacag$

Suffix array is an array of n indices. Thus, it takes O(n log n) bits.

Suffix Position SA[i] Suffix

acacag$ 1 7 $

cacag$ 2 1 acacag$

acag$ 3 => 3 acag$

cag$ 4 Sort 5 ag$

ag$ 5 2 cacag$

g$ 6 4 cag$

$ 7 6 g$

Observation The leaves of a suffix tree is in

suffix array order.

$ a ca

g

cag

cg

c g

7

1 3

5

2 4

6

ag

g$ $

$

$

$$a

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

Linear time construction of suffix array from suffix tree Recall that the suffix tree T of S[1..n] can be

constructed in O(n) time. Then, by “lexical” depth-first traversal of T,

the suffix array of S is obtained. This takes O(n) time.

However, the space used during construction is the same as that for suffix tree! This defeats the purpose of suffix array.

Today, we can build suffix array using O(n) bit space and O(n) time.

range(T,Q) For a pattern Q, its

occurrences in T form a consecutive SA range.

Example: For T= acacag$, ca occurs in SA[5] and SA[6].

Definition: We called range(T,Q)=[st..ed] if

Q is a prefix of every Tj for j=SA[st], SA[st+1], …, SA[ed]

where Tj = j suffix of T = T[j..n]. Example: range(T,ca)=[5..6]

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

1

2

3

4

5

6

7

Find occurrence of query Q in a string S using suffix array Input: (1) the suffix array of a string T

of length n and (2) a query Q of length m

Aim: check if Q occurs in T

Idea: binary search!

Algorithm

Example Consider T = acacag$ Pattern Q = acag

L=1 R=7 M=(L+R)/2=4

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

i

1

2

3

4

5

6

7


L=1 R=7 M=(L+R)/2=4

suffix-SA[M] > Q. Set R=M=4.

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

i

1

2

3

4

5

6

7


L=1 R=4 M=(L+R)/2=2

suffix-SA[M] < Q. Set L=M=2.

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

i

1

2

3

4

5

6

7


L=2 R=4 M=(L+R)/2=3

The pattern Q is found at SA[M]=3.

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

i

1

2

3

4

5

6

7

Can we do better? During each step of binary search,

we need to compare Q with a suffix using O(m) time, which is time consuming.

Can we do better?

We have the following observation. Suppose LCP(Q, suffix-SA[L]) is l and LCP(Q, suffix-SA[R])

is r. Then, LCP(Q, suffix-SA[M]) > min{l,r}.

Below, we describe how to utilize this observation to speedup the computation.

Algorithm


L=1, l=0 R=7, r=0 mlr = min(l,r)=0 M=(L+R)/2=4

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

i

1

2

3

4

5

6

7


L=1, l=0 R=7, r=0 mlr = min(l,r)=0 M=(L+R)/2=4, m=1

The (m+1) char of suffix-SA[M] is g. The (m+1) char of Q is c. So, suffix-SA[M] > Q. Set R=M=4 and r=m=1.

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

i

1

2

3

4

5

6

7



The (m+1) char of suffix-SA[M] is c. The (m+1) char of Q is g. So, suffix-SA[M] < Q. Set L=M=2 and l=m=3.

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

i

1

2

3

4

5

6

7



The pattern Q is found at SA[M]=3.

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

i

1

2

3

4

5

6

7

Time analysis Binary search will perform log n

comparisons Each comparison takes at most

O(m) time In the worst case, O(m log n) time. Myers and Manber report that, in

practice, the time is O(m + log n).

Suffix array and suffix tree We show one example of replacing

suffix tree by suffix array Note that most applications related

to suffix tree can be solved using suffix array with some time blow up!

When space is limited, replacing suffix tree by suffix array is a good choice.

The size is still too big! Why? DNA sequences can be very long!

E.g. Fly: ~100M bases, Human: ~3G bases, Tree: ~9G bases

Storage to store indexing data structure for human genomeSuffix Tree: ~40G bytesSuffix Array: ~13G bytes

Can we further reduce the space?

Solution Grossi, Vitter (STOC2000)

Compressed suffix array (CSA) Ferragine, Manzini (FOCS2000)

FM-index Both of them can be stored in O(n) bit

space For Human Genome

Both CSA and FM-index can be stored within 2G bytes.

FM-index Consider a text T=acacag$ FM-index stores:

A. The string BW=g$ccaaaB. C[x] = total no. of

occurrences of each symbol less than x. E.g. C[a]=1, C[c]=4, C[g]=6, C[t]=7

C. A data-structure occ(x, i) which tells us the number of occurrences of x in BW[1..i] using O(1) time.

SA[i] Suffix T[SA[i]-1]

7 $ g

1 acacag$ $

3 acag$ c

5 ag$ c

2 cacag$ a

4 cag$ a

6 g$ a

Data-structure for answering the occ(x, i) query?

Given the text BW[1..n], we divide BW[1..n] into buckets of size log2n. For each bucket i = 1, …, n/log2n, we store

P[i] = number of x’s in BW[1.. ilog2n]. Each bucket is further subdivided into log n sub-buckets of size log n. For each sub-bucket j of the bucket i, we store

Q[i][j] = number of x’s in the first j sub-buckets.

………BWlog2 n

………

log n

2 n/log2 n3 41

1 2 3 log n

00x0xxx0x0

Data-structure for answering the occ(x, i) query? We also need a lookup table rank(b,k)

b is any string of length (log n)/2 1 k (log n)/2

rank(b,k) = number of x in the first k characters of b.

00…0

xx…x

1 2 3 ………………(log n)/2

x0x..0

……

1 1 2 ………………………

bits )()logloglog(

space Total

bits )log(log sentry takeEach

2

log

2

log2

entries ofNumber

2

log

nonnnO

nO

nn

nn

Space complexity of the occ() data-structure P[1..n/log2n] uses O(n/log n) bits Q[1..n/log2n][1..log n] uses O(n log

log n / log n) bits rank(b,k) uses 2logn/2 (log n/2) =

o(n) bits

In total, we use O(n log log n / log n) bits.

How to compute occ(x,i)?

Suppose log n = 10. To compute occ(x, 327),

The result is P[3]+Q[4][2]+rank(00x0x,5)+rank(xx0x0,2)

Hence, O(1) time to compute occ(x,i).

………BWlog2 n

………

log n

2 n/log2 n3 41

1 2 3 log n

00x0xxx0x0

Size of FM-index Structure A can be store in 2n bits Structure B can be store in O(log n)

bits Structure C can be store in

O(n log log n/log n) bits

In total, the size of FM-index isO(n) bits.

Observation C[x]+occ(x,i) is the

number of suffixes smaller than of xTSA[i].

Example: C[c]=4 occ(c, 6)=2 Number of suffixes smaller

than cT[SA[6]..n]=ccag$ is 6.


7 $ g

1 acacag$ $

3 acag$ c

5 ag$ c

2 cacag$ a

4 cag$ a

6 g$ a

Lemma Suppose range(T,Q) is [st..ed]. Then, range(T,xQ) = [p..q] where

p = C[x] + occ(x,st-1) + 1 q = C[x] + occ(x,ed)

Proof: p = 1+number of suffixes strictly smaller

than xQ. The latter term = number of suffixes smaller than or equal to xTSA[st-1].

q = number of suffixes smaller than or equal to xTSA[ed].

Backward search Given the text T and the FM-index, we want to

determine if Q exists in T.

Algorithm BW_exist(Q[1..m])1. x=Q[m], i=m-1;2. /* find range(T,Q[m]) */

st = C[x]+1, ed = C[x+1];3. while (sted and i>1) {

/* find range(T, Q[i-1..m]) */ x = Q[i-1]; st = C[x] + occ(x, st-1) + 1; ed = C[x] + occ(x, ed); i = i – 1;}

4. if st > ed, then pattern not found else pattern found.

Example T = acacag$ Q = aca

Q[3..3]=a sp=C[a] +1

=1+1=2 ep=C[c]

=4


7 $ g

1 acacag$ $

3 acag$ c

5 ag$ c

2 cacag$ a

4 cag$ a

6 g$ a

g $ c c a a a


Q[2..3]=ca st=C[c]+occ(c,stold-1)+1

=4+0+1=5 ed=C[c]+occ(c,edold)

=4+2=6


7 $ g

1 acacag$ $

3 acag$ c

5 ag$ c

2 cacag$ a

4 cag$ a

6 g$ a

g $ c c a a a


Q[1..3]=aca st=C[a]+occ(a,stold-1)+1

=1+0+1=2 ed=C[a]+occ(a,edold)

=1+2=3


7 $ g

1 acacag$ $

3 acag$ c

5 ag$ c

2 cacag$ a

4 cag$ a

6 g$ a

g $ c c a a a


Q[1..3]=aca st=C[a]+occ(a,stold-1)+1

=1+0+1=2 ed=C[a]+occ(a,edold)

=1+2=3


7 $ g

1 acacag$ $

3 acag$ c

5 ag$ c

2 cacag$ a

4 cag$ a

6 g$ a

g $ c c a a a

Q occurrences in T!

Time complexity of Backward Search To find a pattern Q[1..m]

Step 1, 2, and 4 can be computed in O(1) time.

For step 3, We need to iterate the loop for m-1 times. Each iteration of the loop can be computed

in O(1) time. The loop takes O(m) time.

In total, O(m) time for backward search.

Conclusion Suffix tree is a powerful data-structure

which has a lot of applications in Computational Biology.

Problems: Suffix tree is too big!

Can be solved using CSA and FM-index Suffix tree can only solve exact match

problem! (Most of the biology problems are approximate match!)

Many works have been done on this area! But still not pratical. One of the important area to explore!

1-mismatch problem

1-mismatch problem Index: the suffix tree of a text T[1..n] For any pattern P[1..m], the 1-mismatch

problem finds all occurrences of P in T that has hamming distance at most 1.

Example:P = ACGTT = AACGTGGCCAACTTGGA

Naïve solution Index:

create the suffix tree for T Algorithm for query:

Generate all possible 1-mismatch patterns of P. Find occurrences of every 1-mismatch pattern

Running time: There are ||m possible 1-mismatch patterns. Using suffix tree, it takes O(m) time to find

occurrences of each 1-mismatch pattern In total, O(m2+occ) time where occ is total

number of occurrences.

Any other solutions? Cobbs [CPM 1995]

O(n log n) bit index, O(m2+occ) query time. Amir et al [Journal of Algorithm 2000]

O(n log3 n) bit index, O(m log n loglog n + occ) query time. Buchsbaum et al [ESA 2000]

O(n log2 n) bit index, O(m log log n + occ) query time. Cole et al [SODA 2004]

O(n log2 n) bit index, O(m + log log n + occ) query time. Trinh et al [CPM 2004]

O(n log n) bit index, O(m log n + occ) query time. Lam et al [ISAAC 2005]

O(n sqrt(log n)) bit index, O(m log log n + occ) query time. Chan et al [ESA 2006]

O(n log n) bit index, O(m + log log n + occ) query time.

Today, we have a look of the solution of Trinh et al.

Index Suffix array of T

SA[1..n] Inversed suffix array of T

SA-1[1..n] where SA[SA-1[i]] = i.

Definition: We called range(T,P)=[st..ed] if

P is a prefix of every Tj for j=SA[st], SA[st+1], …, SA[ed]

where Tj = j suffix of T = T[j..n]. Example: For T= acacag$, range(T,ca)=[5..6]

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

i

1

2

3

4

5

6

7

Lemma 1 (Forward search) Assume [st..ed]=range(T,P). We can compute

[st’..ed’]=range(T,Pc) in O(log n) time. Proof: binary search on SA.

Lemma 2 Assume [st1..ed1]=range(T,P1) and [st2..ed2]=range(T,P2). We can compute [st..ed]=range(T,P1P2) in O(log n) time.

Proof: Let the length of P1 be k. Note that TSA[st1],TSA[st1+1],…,TSA[ed1] are lexicographically increasing. Hence, TSA[st1]+k,TSA[st1+1]+k,…,TSA[ed1]+k are lexicographically

increasing. Thus, SA-1[SA[st1]+k] < SA-1[SA[st1+1]+k] <…< SA-1[SA[ed1]+k]. To find st and ed, we need to find

the smallest st such that st2 < SA-1[SA[st]+k]<ed2 and the largest ed such that st2 < SA-1[SA[ed]+k] < ed2.

This can be done by binary search.

Example T=acacag$ P1=a, P2=ca. range(T,P1)=[2..4], range(T,P2)=[5..6] To find range(T,P1P2), we do the following:

Note that TSA[2]<TSA[3]<TSA[4] are beginning with a. So, TSA[2]+1<TSA[3]+1<TSA[4]+1. Note that

SA-1[SA[2]+1]=5, SA-1[SA[3]+1]=6, SA-1[SA[4]+1]=7. Hence, TSA[5]<TSA[6]<TSA[7]. Among the three suffixes, we need to identify suffix beginning

with P2. Since range(T,P2)=[5..6], both TSA[5] and TSA[6] contain P2. As SA-1[SA[2]+1]=5 and SA-1[SA[3]+1]=6,

we have range(T,P1P2) = [2..3].

SA[i] Suffix

7 $

1 acacag$

3 acag$

5 ag$

2 cacag$

4 cag$

6 g$

i

1

2

3

4

5

6

7

Lemma 3 (Backward search) Assume [st..ed]=range(T,P). We can compute

[st’..ed’]=range(T,cP) in O(log n) time. Proof: Let P1=c, P2=P. By Lemma 2,

range(T,P1P2) can be computed in O(log n) time.

Algorithm1. For j=m,m-1,…,1,

By backward search, find range(T,P[j..m]).

2. For j=1,2,…,m, By forward search, find range(T,P[1..j]).

3. Report all occurrences of range(T,P[1..m])4. For j=1,2,…,m,

Let P1=P[1..j-1], P2=P[j+1..m] For every character cP[j],

By forward search, find range(T,P1c) By Lemma 2, find range(T,P1cP2) Report all occurrences of range(T,P1cP2)

Time analysis Ignoring the O(occ) reporting time! Steps 1 and 2 take O(m log n) time. Step 4 tries all possible mismatches.

There are in total ||m mismatches. For each mismatch, it take O(log n) time. So, Step 4 takes O(||m log n) time.

In total, the running time is O(||m log n + occ).

Note that this solution can be generalized to handle k-mismatch or k-difference.

algorithms in bioinformatics: a practical introduction

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