# On line construction of suffix trees 1

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1 (To appear in ALGORITHMICA) On line construction of suffix trees 1 Esko Ukkonen Department of Computer Science, University of Helsinki, P. O. Box 26 (Teollisuuskatu 23), FIN University of Helsinki, Finland Tel.: , fax: Abstract. An on line algorithm is presented for constructing the suffix tree for a given string in time linear in the length of the string. The new algorithm has the desirable property of processing the string symbol by symbol from left to right. It has always the suffix tree for the scanned part of the string ready. The method is developed as a linear time version of a very simple algorithm for (quadratic size) suffix tries. Regardless of its quadratic worst-case this latter algorithm can be a good practical method when the string is not too long. Another variation of this method is shown to give in a natural way the well known algorithms for constructing suffix automata (DAWGs). Key Words. Linear time algorithm, suffix tree, suffix trie, suffix automaton, DAWG. 1 Research supported by the Academy of Finland and by the Alexander von Humboldt Foundation (Germany). 1

2 1. INTRODUCTION A suffix tree is a trie like data structure representing all suffixes of a string. Such trees have a central role in many algorithms on strings, see e.g. [3, 7, 2]. It is quite commonly felt, however, that the linear time suffix tree algorithms presented in the literature are rather difficult to grasp. The main purpose of this paper is to be an attempt in developing an understandable suffix tree construction based on a natural idea that seems to complete our picture of suffix trees in an essential way. The new algorithm has the important property of being on line. It processes the string symbol by symbol from left to right, and has always the suffix tree for the scanned part of the string ready. The algorithm is based on the simple observation that the suffixes of a string T i = t 1 t i can be obtained from the suffixes of string T i 1 = t 1 t i 1 by catenating symbol t i at the end of each suffix of T i 1 and by adding the empty suffix. The suffixes of the whole string T = T n = t 1 t 2 t n can be obtained by first expanding the suffixes of T 0 into the suffixes of T 1 and so on, until the suffixes of T are obtained from the suffixes of T n 1. This is in contrast with the method by Weiner [13] that proceeds right to left and adds the suffixes to the tree in increasing order of their length, starting from the shortest suffix, and with the method by McCreight [9] that adds the suffixes to the tree in the decreasing order of their length. It should be noted, however, that despite of the clear difference in the intuitive view on the problem, our algorithm and McCreight s algorithm are in their final form functionally rather closely related. Our algorithm is best understood as a linear time version of another algorithm from [12] for (quadratic size) suffix tries. The latter very elementary algorithm, which resembles the position tree algorithm in [8], is given in Section 2. Unfortunately, it does not run in linear time it takes time proportional to the size of the suffix trie which can be quadratic. However, a rather transparent modification, which we describe in Section 4, gives our on line, linear time method for suffix trees. This also offers a natural perspective 2

3 which makes the linear time suffix tree construction understandable. We also point out in Section 5 that the suffix trie augmented with the suffix links gives an elementary characterization of the suffix automata (also known as directed acyclic word graphs or DAWGs). This immediately leads to an algorithm for constructing such automata. Fortunately, the resulting method is essentially the same as already given in [4 6]. Again it is felt that our new perspective is very natural and helps understanding the suffix automata constructions. 2. CONSTRUCTING SUFFIX TRIES Let T = t 1 t 2 t n be a string over an alphabet Σ. Each string x such that T = uxv for some (possibly empty) strings u and v is a substring of T, and each string T i = t i t n where 1 i n + 1 is a suffix of T ; in particular, T n+1 = ɛ is the empty suffix. The set of all suffixes of T is denoted σ(t ). The suffix trie of T is a trie representing σ(t ). More formally, we denote the suffix trie of T as ST rie(t ) = (Q { }, root, F, g, f) and define such a trie as an augmented deterministic finite state automaton which has a tree shaped transition graph representing the trie for σ(t ) and which is augmented with the so called suffix function f and auxiliary state. The set Q of the states of ST rie(t ) can be put in a one to one correspondence with the substrings of T. We denote by x the state that corresponds to a substring x. The initial state root corresponds to the empty string ɛ, and the set F of the final states corresponds to σ(t ). The transition function g is defined as g( x, a) = ȳ for all x, ȳ in Q such that y = xa, where a Σ. The suffix function f is defined for each state x Q as follows. Let x root. Then x = ay for some a Σ, and we set f( x) = ȳ. Moreover, f(root) =. Auxiliary state allows us to write the algorithms in the sequel such that an explicit distinction between the empty and the nonempty suffixes 3

4 (or, between root and the other states) can be avoided. State is connected to the trie by g(, a) = root for every a Σ. We leave f( ) undefined. (Note that the transitions from to root are defined consistently with the other transitions: State corresponds to the inverse a 1 of all symbols a Σ. Because a 1 a = ɛ, we can set g(, a) = root as root corresponds to ɛ.) Following [9] we call f(r) the suffix link of state r. The suffix links will be utilized during the construction of a suffix tree; they have many uses also in the applications (e.g. [11, 12]). Automaton STrie(T ) is identical to the Aho Corasick string matching automaton [1] for the key word set {T i 1 i n + 1} (the suffix links are called in [1] the failure transitions.) Fig. 1. Construction of STrie(cacao): state transitions shown in bold arrows, failure transitions in thin arrows. Note: Only the last two layers of suffix links shown explicitly. 4

5 It is easy to construct STrie(T ) on line, in a left to right scan over T as follows. Let T i denote the prefix t 1 t i of T for 0 i n. As intermediate results the construction gives STrie(T i ) for i = 0, 1,..., n. Fig. 1 shows the different phases of constructing STrie(T ) for T = cacao. The key observation explaining how STrie(T i ) is obtained from STrie(T i 1 ) is that the suffixes of T i can be obtained by catenating t i to the end of each suffix of T i 1 and by adding an empty suffix. That is, σ(t i ) = σ(t i 1 )t i {ɛ}. By definition, STrie(T i 1 ) accepts σ(t i 1 ). To make it accept σ(t i ), we must examine the final state set F i 1 of ST rie(t i 1 ). If r F i 1 has not already a t i transition, such a transition from r to a new state (which becomes a new leaf of the trie) is added. The states to which there is an old or new t i transition from some state in F i 1 constitute together with root the final states F i of STrie(T i ). The states r F i 1 that get new transitions can be found using the suffix links as follows. The definition of the suffix function implies that r F i 1 if and only if r = f j (t 1... t i 1 ) for some 0 j i 1. Therefore all states in F i 1 are on the path of suffix links that starts from the deepest state t 1... t i 1 of STrie(T i 1 ) and ends at. We call this important path the boundary path of ST rie(t i 1 ). The boundary path is traversed. If a state z on the boundary path does not have a transition on t i yet, a new state zt i and a new transition g( z, t i ) = zt i are added. This gives updated g. To get updated f, the new states zt i are linked together with new suffix links that form a path starting from state t 1... t i. Obviously, this is the boundary path of ST rie(t i ). The traversal over F i 1 along the boundary path can be stopped immediately when the first state z is found such that state zt i (and hence also transition g( z, t i ) = zt i ) already exists. Let namely zt i already be a state. Then STrie(T i 1 ) has to contain state z t i and transition g(z, t i ) = z t i for all z = f j ( z), j 1. In other words, if zt i is a substring of T i 1 then every suffix of zt i is a substring of T i 1. Note that z always exists because is the 5

6 last state on the boundary path and has a transition for every possible t i. When the traversal is stopped in this way, the procedure will create a new state for every suffix link examined during the traversal. This implies that the whole procedure will take time proportional to the size of the resulting automaton. Summarized, the procedure for building STrie(T i ) from STrie(T i 1 ) is as follows [12]. Here top denotes the state t 1... t i 1. Algorithm 1. r top; while g(r, t i ) is undefined do create new state r and new transition g(r, t i ) = r ; if r top then create new suffix link f(oldr ) = r ; oldr r ; r f(r); create new suffix link f(oldr ) = g(r, t i ); top g(top, t i ). Starting from STrie(ɛ), which consists only of root and and the links between them, and repeating Algorithm 1 for t i = t 1, t 2,..., t n, we obviously get STrie(T ). The algorithm is optimal in the sense that it takes time proportional to the size of its end result STrie(T ). This in turn is proportional to Q, that is, to the number of different substrings of T. Unfortunately, this can be quadratic in T, as is the case for example if T = a n b n. Theorem 1 Suffix trie ST rie(t ) can be constructed in time proportional to the size of ST rie(t ) which, in the worst case, is O( T 2 ). 3. SUFFIX TREES Suffix tree STree(T ) of T is a data structure that represents STrie(T ) in space linear in the length T of T. This is achieved by representing only a subset Q { } of the states of STrie(T ). We call the states in Q { } 6

7 the explicit states. Set Q consists of all branching states (states from which there are at least two transitions) and all leaves (states from which there are no transitions) of STrie(T ). By definition, root is included into the branching states. The other states of STrie(T ) (the states other than root and from which there is exactly one transition) are called implicit states as states of STree(T ); they are not explicitly present in STree(T ). The string w spelled out by the transition path in STrie(T ) between two explicit states s and r is represented in STree(T ) as generalized transition g (s, w) = r. To save space the string w is actually represented as a pair (k, p) of pointers (the left pointer k and the right pointer p) to T such that t k... t p = w. In this way the generalized transition gets form g (s, (k, p)) = r. Such pointers exist because there must be a suffix T i such that the transition path for T i in STrie(T ) goes through s and r. We could select the smallest such i, and let k and p point to the substring of this T i that is spelled out by the transition path from s to r. A transition g (s, (k, p)) = r is called an a transition if t k = a. Each s can have at most one a transition for each a Σ. Transitions g(, a) = root are represented in a similar fashion: Let Σ = {a 1, a 2,..., a m }. Then g(, a j ) = root is represented as g(, ( j, j)) = root for j = 1,..., m. Hence suffix tree STree(T ) has two components: The tree itself and the string T. It is of linear size in T because Q has at most T leaves (there is at most one leaf for each nonempty suffix) and therefore Q has to contain at most T 1 branching states (when T > 1). There can be at most 2 T 2 transitions between the states in Q, each taking a constant space because of using pointers instead of an explicit string. (Here we have assumed the standard RAM model in which a pointer takes constant space.) We again augment the structure with the suffix function f, now defined only for all branching states x root as f ( x) = ȳ where y is a branching state such that x = ay for some a Σ, and f (root) =. Such an f is well defined: If x is a branching state, then also f ( x) is a branching state. These suffix links are explicitly represented. It will sometimes be helpful 7

8 to speak about implicit suffix links, i.e. imaginary suffix links between the implicit states. The suffix tree of T is denoted as STree(T ) = (Q { }, root, g, f ). We refer to an explicit or implicit state r of a suffix tree by a reference pair (s, w) where s is some explicit state that is an ancestor of r and w is the string spelled out by the transitions from s to r in the corresponding suffix trie. A reference pair is canonical if s is the closest ancestor of r (and hence, w is shortest possible). For an explicit r the canonical reference pair obviously is (r, ɛ). Again, we represent string w as a pair (k, p) of pointers such that t k... t p = w. In this way a reference pair (s, w) gets form (s, (k, p)). Pair (s, ɛ) is represented as (s, (p + 1, p)). It is technically convenient to omit the final states in the definition of a suffix tree. When explicit final states are needed in some application, one gets them gratuitously by adding to T an end marking symbol that does not occur elsewhere in T. The leaves of the suffix tree for such a T are in one to one correspondence with the suffixes of T and constitute the set of the final states. Another possibility is to traverse the suffix link path from leaf T to root and make all states on the path explicit; these states are the final states of STree(T ). In many applications of STree(T ), the start location of each suffix is stored with the corresponding state. Such an augmented tree can be used as an index for finding any substring of T. 4. ON LINE CONSTRUCTION OF SUFFIX TREES The algorithm for constructing STree(T ) will be patterned after Algorithm 1. What has to be done is for the most part immediately clear. Fig. 2 shows the phases of constructing STree(cacao); for simplicity, the strings associated with each transition are shown explicitly in the figure. However, to get a linear time algorithm some details need a more careful examination. We first make more precise what Algorithm 1 does. Let s 1 = t 1... t i 1, s 2, s 3,..., s i = root, s i+1 = be the states of STrie(T i 1 ) on the boundary 8

9 path. Let j be the smallest index such that s j is not a leaf, and let j be the smallest index such that s j has a t i transition. As s 1 is a leaf and is a non leaf that has a t i transition, both j and j are well defined and j j. Now the following lemma should be obvious. Fig. 2. Construction of STree(cacao) Lemma 1 Algorithm 1 adds to ST rie(t i 1 ) a t i transition for each of the states s h, 1 h < j, such that for 1 h < j, the new transition expands an old branch of the trie that ends at leaf s h, and for j h < j, the new transition initiates a new branch from s h. Algorithm 1 does not create any other transitions. We call state s j the active point and s j the end point of ST rie(t i 1 ). These states are present, explicitly or implicitly, in ST ree(t i 1 ), too. For example, the active points of the last three trees in Fig. 2 are (root, c), (root, ca), (root, ɛ). Lemma 1 says that Algorithm 1 inserts two different groups of t i transitions into ST rie(t i 1 ): (i) First, the states on the boundary path before the active point s j get a transition. These states are leaves, hence each such transition has to expand 9

10 an existing branch of the trie. (ii) Second, the states from the active point s j to the end point s j, the end point excluded, get a new transition. These states are not leaves, hence each new transition has to initiate a new branch. Let us next interpret this in terms of suffix tree STree(T i 1 ). The first group of transitions that expand an existing branch could be implemented by updating the right pointer of each transition that represents the branch. Let g (s, (k, i 1)) = r be such a transition. The right pointer has to point to the last position i 1 of T i 1. This is because r is a leaf and therefore a path leading to r has to spell out a suffix of T i 1 that does not occur elsewhere in T i 1. Then the updated transition must be g (s, (k, i)) = r. This only makes the string spelled out by the transition longer but does not change the states s and r. Making all such updates would take too much time. Therefore we use the following trick. Any transition of STree(T i 1 ) leading to a leaf is called an open transition. Such a transition is of the form g (s, (k, i 1)) = r where, as stated above, the right pointer has to point to the last position i 1 of T i 1. Therefore it is not necessary to represent the actual value of the right pointer. Instead, open transitions are represented as g (s, (k, )) = r where indicates that this transition is open to grow. In fact, g (s, (k, )) = r represents a branch of any length between state s and the imaginary state r that is in infinity. An explicit updating of the right pointer when t i is inserted into this branch is not needed. Symbols can be replaced by n = T after completing STree(T ). In this way the first group of transitions is implemented without any explicit changes to ST ree(t i 1 ). We have still to describe how to add to STree(T i 1 ) the second group of transitions. These create entirely new branches that start from states s h, j h < j. Finding such states s h needs some care as they need not be explicit states at the moment. They will be found along the boundary path of ST ree(t i 1 ) using reference pairs and suffix links. Let h = j and let (s, w) be the canonical reference pair for s h, i. e., for 10

11 the active point. As s h is on the boundary path of STrie(T i 1 ), w has to be a suffix of T i 1. Hence (s, w) = (s, (k, i 1)) for some k i. We want to create a new branch starting from the state represented by (s, (k, i 1)). However, first we test whether or not (s, (k, i 1)) already refers to the end point s j. If it does, we are done. Otherwise a new branch has to be created. To this end the state s h referred to by (s, (k, i 1)) has to be explicit. If it is not, an explicit state, denoted s h, is created by splitting the transition that contains the corresponding implicit state. Then a t i transition from s h is created. It has to be an open transition g (s h, (i, )) = s h where s h is a new leaf. Moreover, the suffix link f (s h ) is added if s h was created by splitting a transition. Next the construction proceeds to s h+1. As the reference pair for s h was (s, (k, i 1)), the canonical reference pair for s h+1 is canonize(f (s), (k, i 1)) where canonize makes the reference pair canonical by updating the state and the left pointer (note that the right pointer i 1 remains unchanged in canonization). The above operations are then repeated for s h+1, and so on until the end point s j is found. In this way we obtain the procedure update, given below, that transforms STree(T i 1 ) into STree(T i ) by inserting the t i transitions in the second group. The procedure uses procedure canonize mentioned above, and procedure test and split that tests whether or not a given reference pair refers to the end point. If it does not then the procedure creates and returns an explicit state for the reference pair provided that the pair does not already represent an explicit state. Procedure update returns a reference pair for the end point s j (actually only the state and the left pointer of the pair, as the second pointer remains i 1 for all states on the boundary path). 11

12 procedure update(s, (k, i)): (s, (k, i 1)) is the canonical reference pair for the active point; 1. oldr root; (end point, r) test and split(s, (k, i 1), t i ); 2. while not(end point) do 3. create new transition g (r, (i, )) = r where r is a new state; 4. if oldr root then create new suffix link f (oldr) = r; 5. oldr r; 6. (s, k) canonize(f (s), (k, i 1)); 7. (end point, r) test and split(s, (k, i 1), t i ); 8. if oldr root then create new suffix link f (oldr) = s; 9. return (s, k). Procedure test and split tests whether or not a state with canonical reference pair (s, (k, p)) is the end point, that is, a state that in STrie(T i 1 ) would have a t i transition. Symbol t i is given as input parameter t. The test result is returned as the first output parameter. If (s, (k, p)) is not the end point, then state (s, (k, p)) is made explicit (if not already so) by splitting a transition. The explicit state is returned as the second output parameter. procedure test and split(s, (k, p), t): 1. if k p then 2. let g (s, (k, p )) = s be the t k transition from s; 3. if t = t k +p k+1 then return(true, s) 4. else 5. replace the t k transition above by transitions g (s, (k, k + p k)) = r and g (r, (k + p k + 1, p )) = s where r is a new state; 6. return(false, r) 7. else 8. if there is no t transition from s then return(false, s) 9. else return(true, s). 12

13 This procedure benefits from that (s, (k, p)) is canonical: The answer to the end point test can be found in constant time by considering only one transition from s. Procedure canonize is as follows. Given a reference pair (s, (k, p)) for some state r, it finds and returns state s and left link k such that (s, (k, p)) is the canonical reference pair for r. State s is the closest explicit ancestor of r (or r itself if r is explicit). Therefore the string that leads from s to r must be a suffix of the string t k... t p that leads from s to r. Hence the right link p does not change but the left link k can become k, k k. procedure canonize(s, (k, p)): 1. if p < k then return (s, k) 2. else 3. find the t k transition g (s, (k, p )) = s from s; 4. while p k p k do 5. k k + p k + 1; 6. s s ; 7. if k p then find the t k transition g (s, (k, p )) = s from s; 8. return (s, k). To be able to continue the construction for the next text symbol t i+1, the active point of STree(T i ) has to be found. To this end, note first that s j is the active point of ST ree(t i 1 ) if and only if s j = t j t i 1 where t j t i 1 is the longest suffix of T i 1 that occurs at least twice in T i 1. Second, note that s j is the end point of ST ree(t i 1 ) if and only if s j = t j t i 1 where t j t i 1 is the longest suffix of T i 1 such that t j t i 1 t i is a substring of T i 1. But this means that if s j is the end point of ST ree(t i 1 ) then t j t i 1 t i is the longest suffix of T i that occurs at least twice in T i, that is, then state g(s j, t i ) is the active point of ST ree(t i ). We have shown the following result. Lemma 2 Let (s, (k, i 1)) be a reference pair of the end point s j of ST ree(t i 1 ). Then (s, (k, i)) is a reference pair of the active point of ST ree(t i ). 13

14 The overall algorithm for constructing STree(T ) is finally as follows. String T is processed symbol by symbol, in one left-to-right scan. Writing Σ = {t 1,..., t m } makes it possible to present the transitions from in the same way as the other transitions. Algorithm 2. Construction of STree(T ) for string T = t 1 t 2... in alphabet Σ = {t 1,..., t m }; is the end marker not appearing elsewhere in T. 1. create states root and ; 2. for j 1,..., m do create transition g (, ( j, j)) = root; 3. create suffix link f (root) = ; 4. s root; k 1; i 0; 5. while t i+1 do 6. i i + 1; 7. (s, k) update(s, (k, i)); 8. (s, k) canonize(s, (k, i)). Steps 7 8 are based on Lemma 2: After step 7 pair (s, (k, i 1)) refers to the end point of ST ree(t i 1 ), and hence, (s, (k, i)) refers to the active point of ST ree(t i ). Theorem 2 Algorithm 2 constructs the suffix tree ST ree(t ) for a string T = t 1... t n on line in time O(n). Proof. The algorithm constructs STree(T ) through intermediate trees STree(T 0 ), STree(T 1 ),..., STree(T n ) = STree(T ). It is on line as to construct STree(T i ) it only needs access to the first i symbols of T. For the running time analysis we divide the time requirement into two components, both turn out to be O(n). The first component consists of the total time for procedure canonize. The second component consists of the rest: The time for repeatedly traversing the suffix link path from the present active point to the end point and creating the new branches by update and then finding the next active point by taking a transition from the end point 14

15 (step 8 of Alg. 2). We call the states (reference pairs) on these paths the visited states. The second component takes time proportional to the total number of the visited states, because the operations at each such state (create an explicit state and a new branch, follow an explicit or implicit suffix link, test for the end point) at each such state can be implemented in constant time as canonize is excluded. (To be precise, this also requires that Σ is bounded independently of n.) Let r i be the active point of STree(T i ) for 0 i n. The visited states between r i 1 and r i are on a path that consists of some suffix links and one t i transition. Taking a suffix link decreases the depth (the length of the string spelled out on the transition path from root) of the current state by one, and taking a t i transition increases it by one. The number of the visited states (including r i 1, excluding r i ) on the path is therefore depth(r i 1 ) depth(r i ) + 2, and their total number is n i=1 (depth(r i 1 ) depth(r i ) + 2) = depth(r 0 ) depth(r n ) + 2n 2n. This implies the second time component is O(n). The time spent by each execution of canonize has an upper bound of the form a + bq where a and b are constants and q is the number of executions of the body of the loop in steps 5 7 of canonize. The total time spent by canonize has therefore a bound that is proportional to the sum of the number of the calls of canonize and the total number of the executions of the body of the loop in all calls. There are O(n) calls as there is one call for each visited state (either in step 6 of update or directly in step 8 of Alg. 2.). Each execution of the body deletes a nonempty string from the left end of string w = t k... t p represented by the pointers in reference pair (s, (k, p)). String w can grow during the whole process only in step 8 of Alg. 2 which catenates t i for i = 1,..., n to the right end of w. Hence a non empty deletion is possible at most n times. The total time for the body of the loop is therefore O(n), and altogether canonize or our first component needs time O(n). 15

16 Remark 1. (due to J. Kärkkäinen) In its final form our algorithm is a rather close relative of McCreight s method [9]. The principal technical difference seems to be, that each execution of the body of the main loop of our Algorithm 2 consumes one text symbol t i whereas each execution of the body of the main loop of McCreight s algorithm traverses one suffix link and consumes zero or more text symbols. Remark 2. It is not hard to generalize Algorithm 2 for the following dynamic version of the suffix tree problem (c.f. the adaptive dictionary matching problem of [2]): Maintain a generalized linear size suffix tree representing all suffixes of strings T i in set {T 1,..., T k } under operations that insert or delete a string T i. The resulting algorithm will make such updates in time O( T i ). 5. CONSTRUCTING SUFFIX AUTOMATA The suffix automaton SA(T ) of a string T = t 1... t n is the minimal DFA that accepts all the suffixes of T. As our STrie(T ) is a DFA for the suffixes of T, SA(T ) could be obtained by minimizing STrie(T ) in standard way. Minimization works by combining the equivalent states, i. e., states from which STrie(T ) accepts the same set of strings. Using the suffix links we will obtain a natural characterization of the equivalent states as follows. A state s of STrie(T ) is called essential if there is at least two different suffix links pointing to s or s = t 1 t k for some k. Theorem 3 Let s and r be two states of ST rie(t ). The set of strings accepted from s is equal to the set of strings accepted from r if and only if the suffix link path that starts from s contains r (the path from r contains s) and the subpath from s to r (from r to s) does not contain any other essential states than possibly s (r). Proof. The theorem is implied by the following observations. The set of strings accepted from some state of STrie(T ) is a subset of the suffixes of T and therefore each accepted string is of different length. 16

17 A string of length i is accepted from a state s of STrie(T ) if and only if the suffix link path that starts from state t 1 t n i contains s. The suffix links form a tree that is directed to its root root. This suggests a method for constructing SA(T ) with a modified Algorithm 1. The new feature is that the construction should create a new state only if the state is essential. An unessential state s is merged with the first essential state that is before s on the suffix link path through s. This is correct as, by Theorem 3, the states are equivalent. As there are O( T ) essential states, the resulting algorithm can be made to work in linear time. The algorithm turns out to be similar to the algorithms in [4 6]. We therefore omit the details. Acknowledgements. J. Kärkkäinen pointed out some inaccuracies in the earlier version [10] of this work. The author is also indebted to E. Sutinen, D. Wood, and, in particular, S. Kurtz and G. A. Stephen for several useful comments. References 1. A. Aho and M. Corasick, Efficient string matching: An aid to bibliographic search, Comm. ACM 18 (1975), A. Amir and M. Farach, Adaptive dictionary matching, in Proc. 32nd IEEE Ann. Symp. on Foundations of Computer Science, 1991, pp A. Apostolico, The myriad virtues of subword trees, in Combinatorial Algorithms on Words (A. Apostolico and Z. Galil, eds.), Springer Verlag, 1985, pp A. Blumer & al., The smallest automaton recognizing the subwords of a text, Theor. Comp. Sci. 40 (1985),

18 5. M. Crochemore, Transducers and repetitions, Theor. Comp. Sci. 45 (1986), M. Crochemore, String matching with constraints, in Mathematical Foundations of Computer Science 1988 (M.P. Chytil, L. Janiga and V. Koubek, eds.), Lect. Notes in Computer Science, vol. 324, Springer Verlag, 1988, pp Z. Galil and R. Giancarlo, Data structures and algorithms for approximate string matching, J. Complexity 4 (1988), M. Kempf, R. Bayer and U. Güntzer, Time optimal left to right construction of position trees, Acta Informatica 24 (1987), E. McCreight, A space economical suffix tree construction algorithm, Journal of the ACM 23 (1976), E. Ukkonen, Constructing suffix trees on line in linear time, in Algorithms, Software, Architecture. Information Processing 92, vol. I (J. van Leeuwen, ed.), Elsevier, 1992, pp E. Ukkonen, Approximate string matching over suffix trees, in Combinatorial Pattern Matching, CPM 93 (A. Apostolico, M. Crochemore, Z. Galil, and U. Manber, eds.), Lect. Notes in Computer Science, vol. 684, Springer Verlag, 1993, pp E. Ukkonen and D. Wood, Approximate string matching with suffix automata, Algorithmica 10 (1993), P. Weiner, Linear pattern matching algorithms, in IEEE 14th Ann. Symp. on Switching and Automata Theory, 1973, pp

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