1 Introduction
The BurrowsWheeler transform (BWT) [6] is one of the most favored options both for (a) compressing and (b) indexing data sets. On the one hand, compression programs like bzip2 apply the BWT to achieve high compression rates. For that, they leverage the effect that the BWT built on repetitive data tends to have long character runs, which can be compressed by runlength compression, i.e., representing a substring of a’s by the tuple . On the other hand, selfindexing data structures like the FMindex [11] enhance the BWT to a fulltext selfindex. A combined approach of both compression and indexing is the runlength compressed FMindex [21], representing a BWT with character runs, i.e., maximal repetitions of a character, runlength compressed in bits. This representation can be computed directly in runlength compressed space thanks to Policriti and Prezza [30]. The BWT and its runlength compressed representation have been intensively studied during the past decades (e.g., [12, 1, 14] and the references therein). Contrary to that, a variant, called the bijective BWT (BBWT) [16], is far from being wellstudied despite its mathematically appealing characteristics^{1}^{1}1The BBWT is a bijection between strings without the need of an artificial delimiter needed, e.g., to invert the BWT.. As a matter of fact, we are only aware of one index data structure based on the BBWT [3] and of two nontrivial construction algorithms [5, 2] of the (uncompressed) BBWT, both with the need of additional data structures.
In this article, we shed more light on the connection between the BWT and the BBWT by quadratic time inplace conversion algorithms in Sect. 5 constructing the BWT from the BBWT, or vice versa. We can also perform these conversions in the runlength compressed setting in time with space linear to the number of the character runs (cf. Sect. 4 and Sect. 4), where is the sum of character runs in the BWT and the BBWT.
2 Related Work
Given a text of length , the BWT of is the string obtained by assigning to the character preceding the th lexicographically smallest suffix of (or the last character of if this suffix is the text itself). By this definition, we can construct the BWT with any suffix array [22] construction algorithm. However, storing the suffix array inherently needs bits of space. Crochemore et al. [9] tackled this space problem with an inplace algorithm constructing the BWT in online on the reversed text by simulating queries on a dynamic wavelet tree [17] that would be built on the (growing) BWT. They also gave an algorithm for restoring the text inplace in time.
In the runlength compressed setting, Policriti and Prezza [30] can compute the runlength compressed BWT having character runs in time while using bits of space. They additionally presented an adaption of the wavelet tree on runlength compressed texts, yielding a representation using bits of space with query and update time. Finally, practical improvements of the runlength compressed BWT construction were considered by Ohno et al. [29].
The BBWT is the string obtained by assigning to the last character of the th smallest string in the list of all conjugates of the factors of the Lyndon factorization sorted with respect to the order [23, Def. 4]. Bannai et al. [2] recently revealed a connection between the bijective BWT and suffix sorting by presenting an time BBWT construction algorithm based on SAIS [28]. With dynamic data structures like a dynamic wavelet tree [27], Bonomo et al. [5] could devise an algorithm computing the BBWT in time. With nearly the same techniques, Mantaci et al. [24] presented an algorithm computing the BWT (and simultaneously the suffix array if needed) from the Lyndon factorization. All these construction algorithms need however data structures taking bits of space. However, the latter two (i.e., [5] and [24]) can work inplace by simulating the mapping (cf. Sects. 3.5 and 3.4), which we focus on in Sect. 5.1.
3 Preliminaries
Our computational model is the word RAM model with word size . Accessing a word costs time. An algorithm is called inplace if it uses, besides a rewriteable input, only bits of working space. We write for an interval of natural numbers.
3.1 Strings
Let denote an integer alphabet of size with . We call an element a string. Its length is denoted by . Given an integer , we access the th character of with . Concatenating a string times is abbreviated by . A string is called primitive if there is no string with for an integer with .
When is represented by the concatenation of , i.e., , then , and are called a prefix, substring and suffix of , respectively; the prefix , substring , or suffix is called proper if , , or , respectively. For two integers with , let denote the substring of that begins at position and ends at position in . If , then is the empty string. In particular, the suffix starting at position of is called the th suffix of , and denoted with . An occurrence of a substring in is treated as a subinterval of such that . The longest common prefix (LCP) of two strings and is the longest string that is a prefix of both and .
Orders on Strings.
We denote the lexicographic order with . Given two strings and , if is a prefix of or there exists an integer with such that and . Next we define the order of strings, which is based on the lexicographic order of infinite strings: We write if the infinite concatenation is lexicographically smaller than . For instance, but .
Rank and Select Queries.
Given a string , a character , and an integer , the rank query counts the occurrences of c in , and the select query gives the position of the th c in . We stipulate that . A wavelet tree is a data structure supporting rank and select queries.
3.2 Lyndon Words
Given a string , its th conjugate is defined as for an integer . We say that and all of its conjugates belong to the conjugate class . If a conjugate class contains exactly one conjugate that is lexicographically smaller than all other conjugates, then this conjugate is called a Lyndon word [20]. Equivalently, a string is said to be a Lyndon word if and only if for every proper suffix of [10, Prop. 1.2].
The Lyndon factorization [8] of is the factorization of into a sequence of lexicographically nonincreasing Lyndon words , where (a) each is a Lyndon word for , and (b) for each . Each Lyndon word is called a Lyndon factor.
[[10, Algo. 2.1]] Given a string of length , there is an algorithm that outputs the Lyndon factors one by one in increasing order in total time while keeping only a constant number of pointers to positions in that (a) can move one position forward at one time or (b) can be set to the position of another pointer.
Proof.
The algorithm of Duval uses three variables , , and (cf. LABEL:algoDuval in the appendix) pointing to text positions. is the ending position of the previously computed Lyndon factor (or zero at the beginning). On each step, is incremented by one, while is either incremented by one or reset to , as long as is a prefix of a Lyndon word starting at . If is no longer such a prefix, then is either a Lyndon factor or a repetition of Lyndon factors, each of length . In total, we visit at most characters by incrementing the text positions , , and . ∎
For what follows, we fix a string over an alphabet with size . We use the string as our running example. Its Lyndon factors are , , , and .
3.3 BurrowsWheeler Transforms
We denote the bijective BWT of by , where is the last character of the th string in the list storing the conjugates of all Lyndon factors of sorted with respect to the order. A property of used in this paper as a starting point for an inversion algorithm is the following:
[[5, Lemma 15]] .
Proof.
There is no conjugate of a Lyndon factor that is smaller than the smallest Lyndon factor since for every and every . Therefore, is the smallest string among all conjugates of all Lyndon factors. Hence, is the last character of , which is . ∎
The BWT of , called in the following , is the BBWT of for a delimiter smaller than all other characters in (cf. [15, Lemma 12] since is a Lyndon word). Originally, the BWT is defined by reading the last characters of all cyclic rotations of (without $) sorted lexicographically [6]. Here, we call the resulting string . is equivalent to if contains the aforementioned unique delimiter $. We further write (and analogously or ) to denote the BWT of for a string .
Since (and analogously or ) is a permutation of , it is natural to identify each entry of with a text position: By construction , where is the th lexicographically smallest suffix, i.e., , where is the suffix array of . A similar relation is given between and the circular suffix array [19, 2], which is uniquely defined up to positions of equal Lyndon factors. Figure 1 gives an example for all three variants. In what follows, we review means to simulate a linear traversal of the text in forward or backward manner by , and then translate this result to .
3.4 Backward and Forward Steps
Having the location of in , we can compute (i.e., for ) and (i.e., for ) by rank and select queries. To move from to , which we call a forward step, we can use the mapping:
(1) 
where is the th lexicographically smallest character in . To move from to , we can use the backward step of the FMindex [11], which is also called mapping, and is defined as follows:
(2) 
where is the number of occurrences of those characters in that are smaller than (for each character ). We observe from the second equation of (2) that there is no need for when having . This is important, as we can compute in time only having available. Hence, we can compute in time inplace. However, the same trick does not work with . To lookup , we can use the selection algorithm of Chan et al. [7] using and bits as working space (the algorithm restores after execution) to compute an entry of in time.
In summary, we can compute both and inplace in time. The algorithm of Crochemore et al. [9, Thm. 2] inverting inplace in time uses the result of Munro and Raman [26] computing in time for a constant in the comparison model. As noted by Chan et al. [7, Sect. 1], the time bound for the inversion can be improved to time in the RAM model under the assumption that is rewritable.
If we allow more space, it is still advantageous to favor storing instead of if because storing and in their plain forms take bits and bits, respectively. To compute , we can also compute without by endowing with a predecessor data structure (which we do in Sect. 4.3).
Finally, we also need and on for our conversion algorithms. We can define and similarly for with the following peculiarity:
3.5 Steps in the Bijective BWT
The major difference to the BWT is that the LF mapping of the BBWT can contain multiple cycles, meaning that (or ) recursively applied to a position would result in searching circular (more precisely, the search stays within the same Lyndon factor). This is because is the extended BWT [23, Thm. 20 and Remark 12] applied to the multiset of Lyndon factors
. This fact was exploited for circular pattern matching
[19], but is not of interest here.Instead, we follow the analysis of the socalled rewindings [3, Sect. 3]: Remembering that we store the last character of all conjugates of all Lyndon factors in , we observe that the entries in representing the Lyndon factors (i.e., the last characters of the Lyndon factors) are in sorted order (starting with and ending with ). That is because the lexicographic order and the order are the same for Lyndon words [5, Thm. 8]. Applying the backward step at such an entry results in a rewinding, i.e., we can move from the beginning of a Lyndon factor (represented by in ) to the end of (represented by in ) with one backward step. We use this property with Sect. 3.3 in the following sections to read the Lyndon factors from individually in the order .
4 RunLength Compressed Conversions
We now consider and represented as runlength compressed strings taking and bits of space, where and are the number of character runs in and , respectively. For , the goal of this section is the following:
We can convert to in time using bits as working space, or vice versa.
To prove this theorem, we need a data structure that works in the runlength compressed space while supporting rank and select queries as well as updates more efficiently than the time inplace approach described in Sects. 3.5 and 3.4:
4.1 Runlength Compressed Wavelet Trees
Given a runlength compressed string of uncompressed length with character runs, there is an bits representation of that supports access, rank, select, insertions, and deletions in time [30, Lemma 1]. It consists of (1) a dynamic wavelet tree maintaining the starting characters of each character run and (2) a dynamic Fenwick tree maintaining the lengths of the runs. It can be accelerated to time by using the following representations:

The dynamic wavelet tree of Navarro and Nekrich [27] on a text of length uses bits, and supports both updates and queries in time.

The dynamic Fenwick tree of Bille et al [4, Thm. 2] on bit numbers uses bits, and supports both updates and queries in constant time if updates are restricted to be in/decremental.
The obtained time complexity of this data structure directly improves the construction of : [[30, Thm. 2]] We can construct the in bits of space online on the reversed text in time.
In the runlength compressed wavelet tree representation, and support an update operation and a backward step in time with . This helps us to devise the following two conversions:
4.2 From to
We aim for directly outputting the characters of in reversed order since we can then use the algorithm of Sect. 4.1 building online on the reversed text. We start with the first entry of (corresponding to the last Lyndon factor , i.e., storing according to Sect. 3.3) and do a backward step until we come back at this first entry (i.e., we have visited all characters of ). During that search, we copy the read characters to and mark in an array of length at entry how often we visited the th character run of . Finally, we remove the read cycle of by decreasing the run lengths of by the numbers stored in . By doing so, we remove the last Lyndon factor from and consequently know that the currently first entry of must correspond to . This means that we can apply the algorithm recursively on the remaining to extract and delete the Lyndon factors in reversed order while building in the meantime. By removing , is still a valid BBWT since becomes the BBWT of whose Lyndon factors are the same as of (but without ). Note that it is also possible to build in forward order, i.e., computing for increasing by applying the algorithm of Mantaci et al. [24, Fig. 1] while omitting the suffix array construction.
4.3 From to
To build , we need to be aware of the Lyndon factors of , which we compute with Sect. 3.2 by simulating a forward scan on with on . To this end, we store the entries of the array in a Fusion tree [13] using bits and supporting predecessor search in time.^{2}^{2}2We assume that the alphabet is effective, i.e., that each character of appears at least once in . Otherwise, assume that uses characters. Then we build the static dictionary of Hagerup [18] in time, supporting access to a character in time, assigning each of the characters an integer from . We further map to the alphabet , which can be done in time by using space for a lineartime integer sorting algorithm. This time complexity also covers a forward step in by simulating with the Fusion tree on . Hence, this fusion tree allows us to apply Sect. 3.2 computing the Lyndon factorization of with a multiplicative time penalty since this algorithm only needs to perform forward traversals. The starting point of such a traversal is the position with because returns the first character of . Whenever we detect a Lyndon factor (starting with ), we copy this factor to our dynamic . For that, we always maintain the first and the last position of in memory. Having the last position of , we perform backward steps on until returning at the first position of to read the characters of in reversed order. Then we continue with the algorithm of Sect. 3.2 at the position after (for recursing on ). Inserting a Lyndon factor into works exactly as sketched by Bonomo et al. [5, Thm. 17] or in LABEL:algoConstructBBWT in the appendix (we review this algorithm in detail in Sect. 5.1).
5 InPlace Conversions
We finally present our inplace conversions that work in quadratic time by computing or in time having only stored either , , or . We note that the constructions from the text also work in the comparison model, while inverting a transform or converting two different transforms have a multiplicative time penalty as the fastest option to access in the comparison model uses time for a constant [26]. We start with the construction and inversion of (Sects. 5.2 and 5.1), where we show (a) that we can construct from the text in the same manner as Bonomo et al. [5] construct , and (b) that the latter construction works also inplace. Next, we show in Sect. 5.3 how to invert with the BWT inversion algorithm of Crochemore et al. [9, Fig. 3], which allows us to also convert to with the BWT construction algorithm of the same paper [9, Fig. 2]. Finally, we show a conversion from to in Sect. 5.4. An overview is given in Table 1.
5.1 Constructing and
We can compute and from with the algorithm of Bonomo et al. [5] computing the extended BWT [23]. The extended BWT is the BWT defined on a set of primitive strings. As stated in Sect. 3.5, the extended BWT coincides with if this set of primitive strings is the set of Lyndon factors of [5, Thm. 14]. We briefly describe the algorithm of Bonomo et al. [5] for computing the BBWT (cf. Fig. 2 and LABEL:algoConstructBBWT in the appendix): For each Lyndon factor (starting with ), prepend to . To insert the remaining characters of the factor , let be the position of the currently inserted character. Then perform, for each down to , a backward step , and insert at (cf. LABEL:algoConstructBBWT in the appendix). To understand why this computes , we observe that the last character of the most recently inserted Lyndon factor is always the first character in according to Sect. 3.3. By recursively inserting the preceding character at the place returned by a backward step, we precisely insert this character at the position where we would expect it (another backward step from the same position would then return the inserted character). Using only backward steps and insertions, this algorithm works inplace in time by simulating as described in Sect. 3.4.
Consequently, we can build if is a Lyndon word since in this case and coincide [15, Lemma 12]. That is because sorting the suffixes of is equivalent to sorting the conjugates of (if is a Lyndon word, then its Lyndon factorization consists only of itself).
It is easy to generalize this to work for a general string . First, if is primitive, then we compute its socalled Lyndon conjugate, i.e., a conjugate of that is a Lyndon word. (The Lyndon conjugate of is uniquely defined if is primitive.) We can find the Lyndon conjugate of in time with the following two lemmata:
[ [10, Prop. 1.3] ] Given two Lyndon words and , is a Lyndon word if .
Given a primitive string , we can find its Lyndon conjugate in time with bits of space.
Proof.
Let be the Lyndon conjugate of for . Since is identical to , we are done by running the algorithm of Bonomo et al. [5] on . Finally, if is not primitive, then there is a primitive string such that for an integer . We can compute with the above considerations. For obtaining , according to [25, Prop. 2], we only need to make each character in to a character run of length , i.e., if , we append to for increasing (cf. [15, Thm. 13]). Checking whether is primitive can be done in time by checking for each pair of positions their longest common prefix. We summarized these steps in the pseudo code of LABEL:algoConstructBWTC in the appendix.
5.2 Inverting
To invert , we use the techniques of Crochemore et al. [9, Fig. 3] inverting inplace in time. An invariant is that the BWT entry, whose mapping corresponds to the next character to output, is marked with a unique delimiter $. Given that , the algorithm outputs , sets , removes , and recurses until $ is the last character remaining in . By doing so, it restores the text in text order.
To adapt this algorithm for inverting , we additionally need a pointer storing the first symbol of the text (since there is no unique delimiter such as $ in general). Given that points to , we set and subsequently output . From now on, the algorithm works exactly as [9, Fig. 3] if we set after outputting (cf. LABEL:algoInvertBWTC in the appendix). More involving is inverting or converting to , which we tackle next.
5.3 Inverting
Similarly to Sect. 4.2, we read the Lyndon factors from in the order , and move each read Lyndon factor directly to a text buffer such that while reading the last Lyndon factor for an from , we move the characters of to , producing and . This allows us to recurse by reading always the last Lyndon factor stored in .
Here, we want to apply the inversion algorithm for described in Sect. 5.3. For adapting this algorithm to work with , it suffices to insert $ at (cf. Fig. 3). By doing so, we add $ to the cycle of the currently last Lyndon factor stored in , i.e., we enlarge the Lyndon factor to . That is because (a) corresponds to the last character of (cf. Sect. 3.3), and after inserting $, , hence (a forward step on the last character of gives $) and gives the position in corresponding to . Moreover, inserting $ makes the BBWT of , where is the last Lyndon factor of . We now use the property that is a Lyndon word for each , allowing us to perform the inversion steps of Crochemore et al. [9, Fig. 3] on . By doing so, we can remove the entry of corresponding to for increasing and prepend the extracted characters to the text buffer storing within our working space while keeping a valid BBWT.
Instead of inverting , we can convert to inplace by running the inplace BWT construction algorithm of Crochemore et al. [9, Fig. 2] on the text buffer after the extraction of each Lyndon factor. Unfortunately, this works not characterwise, but needs a Lyndon factor to be fully extracted before inserting its characters into . Interestingly, for the other direction (from to ), we can propose a different kind of conversion that works directly on without decoding it.
5.4 From to on the Fly
Like in Sect. 4.3, we process the Lyndon factors of individually to compute by scanning in text order to simulate Sect. 3.2. Suppose that we have built on with $ being the th Lyndon factor of , and suppose that we have detected the first Lyndon factor . Let f denote the last character of .^{3}^{3}3f stands for final character. Further let and be the position of the last character of and the last character of , respectively, such that and . Let such that and if or otherwise. Since and are Lyndon factors, . Consequently, the suffix (the context of ) is lexicographically smaller than the suffix (the context of ), i.e., . Figure 4 gives an overview of the introduced setting.
Our aim is to change such that a forward or backward step within the characters belonging to always results in a cycle. Informally, we want to cut out of , which additionally allows us to recursively continue with the mapping to find the end of the next Lyndon factor .^{4}^{4}4As a matter of fact, if we now want to restore the text with the modified by , we would only produce . For that, we exchange with (cf. Fig. 5). Then the character (i.e., the first character of ) becomes the next character of $ in terms of the forward step (), while a backwards search on the first character of yields ’s last character ( returns , but now ). This is sufficient as long as for every . Otherwise, it can happen that we change the mapping from the th f of to the th f of (or vice versa) unintentionally. In such a case, we swap some entries in within the f interval of . In detail, we conduct the exchange ( with ), but continue with swapping and unless becomes that f that corresponds to for increasing starting with until or . This may not be sufficient if the characters we swap are identical (cf. Fig. 6). In such a case, we recurse on the interval of , see also LABEL:algoBWTToBBWT in the appendix.
Instead of checking whether we have created a cycle after each swap, we want to compute the exact number of swaps needed for this task. For that we note that exchanging with decrements the values of for every by one. In particular, changes for those f’s in that are between and . Hence, the number of swaps is the number of positions with . The swaps are performed within the range starting with and covering all positions with and since covers all entries whose mapping has changed. However, if starts with a character run of (or of if )^{5}^{5}5For , , and hence, was $ but now is ., swapping the identical characters does not change , and therefore has no effect of changing . Instead, we search the end of this run within to swap the first entry below this run with the first entry of this run, and recurse on swapping entry with entries below of it.
Correctness.
To see why the swaps restore the LF mapping for and the remaining part of the text , we examine those substrings of that we might no longer find with the LF mapping after exchanging with .
In detail, we examine each substring with that is represented in (before changing it) with , and . Due to the LFmapping, , meaning that is the th f in , which stores f’s. After exchanging with , becomes for with . However, for all , did not change. Hence, we only have to focus on the range .
First, suppose that . If we swap with , then is still , but becomes such that we have fixed the substring . This also works in a more general setting: If for every , we can perform swaps like above for all entries in
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