---
title: "Ontology-Mediated Querying with the Description Logic EL: Trichotomy and Linear Datalog Rewritability"
authors:
  - "Carsten Lutz"
  - "Leif Sabellek"
doi: "10.24963/ijcai.2017/164"
url: "https://ijcai.org/proceedings/2017/164"
pdf: "https://ijcai.org/proceedings/2017/0164.pdf"
venue: "Proceedings of the Twenty-Sixth International Joint Conference on Artificial Intelligence"
year: 2017
pages: "1181-1187"
track: "Main track"
keywords:
  - "Multidisciplinary Topics and Applications: Database Systems"
  - "Knowledge Representation, Reasoning, and Logic: Description Logics and Ontologies"
---

# Ontology-Mediated Querying with the Description Logic EL: Trichotomy and Linear Datalog Rewritability

Carsten Lutz, Leif Sabellek

*Proceedings of the Twenty-Sixth International Joint Conference on Artificial Intelligence*, pages 1181-1187. DOI: [10.24963/ijcai.2017/164](https://doi.org/10.24963/ijcai.2017/164)

## Abstract

We consider ontology-mediated queries (OMQs) based on an EL
  ontology and an atomic query (AQ), provide an ultimately
  fine-grained analysis of data complexity and study rewritability
  into linear Datalog-aiming to capture linear recursion in SQL.
  Our main results are that every such OMQ is in AC0, 
  NL-complete or PTime-complete, and that containment in NL
  coincides with rewritability into linear Datalog (whereas
  containment in AC0 coincides with rewritability into first-order
  logic). We establish natural characterizations of the three cases,
  show that deciding linear Datalog rewritability (as well as the
  mentioned complexities) is ExpTime-complete, give a way to
  construct linear Datalog rewritings when they exist, and prove that
  there is no constant bound on the arity of IDB relations in linear
  Datalog rewritings.

## Full text

## 1 Introduction

An important application of ontologies is to enrich data with a semantics and with domain knowledge while also providing additional vocabulary for query formulation [Calvanese et al., 2009; Kontchakov et al., 2013; Bienvenu et al., 2014; Bienvenu and Ortiz, 2015]. The combination of a traditional database query and an ontology can be viewed as a compound query, commonly referred to as an ontology-mediated query (OMQ). Substantial research efforts have been invested into studying OMQs based on description logic (DL) ontologies, with two dominating topics being the data complexity of OMQs [Hustadt et al., 2005; Krisnadhi and Lutz, 2007; Rosati, 2007; Calvanese et al., 2013] and their rewritability into more standard database query languages such as SQL (which in this context is often equated with first-order logic) and into Datalog [P´erez-Urbina et al., 2010; Eiter et al., 2012; Bienvenu et al., 2013; 2014; Kaminski et al., 2014; Trivela et al., 2015; Feier et al., 2017]. While the former topic aims to understand the feasibility of OMQs from a theoretical angle, the latter is inspired by rather practical concerns: since most database systems are unaware of ontologies, rewriting OMQs into standard query languages provides an important avenue for implementing OMQ execution in practical applications; however, a major challenge emerges from the fact that the desired rewritings are typically not guaranteed to always exist, though they often do exist in practically relevant cases. Both topics are thoroughly intertwined since rewritability into firstorder logic (FO) is closely related to AC0 data complexity while rewritability into Datalog is closely related to PTIME data complexity.

Modern DLs can roughly be divided into two families: 'expressive DLs' such as ALC and SHIQ which typically have CONP data complexity and where rewritability is guaranteed neither into FO nor into Datalog [Bienvenu et al., 2014; Trivela et al., 2015; Feier et al., 2017], and 'Horn DLs' such as EL and Horn-SHIQ which typically have PTIME data complexity and where rewritability into Datalog is guaranteed, but FO-rewritability is not [Bienvenu et al., 2013; Hansen et al., 2015; Bienvenu et al., 2016] (with the notable exception of DL-Lite [Calvanese et al., 2009]). In this paper, we consider the OMQ language (EL, AQ) where the ontology is formulated in the Horn description logic EL and where the actual queries are atomic queries (AQs) of the form A(x), studying data complexity, rewritability, and their relations. Our actual contribution is two-fold.

First, we carry out an ultimately fine-grained analysis of data complexity. In fact, we establish a trichotomy, showing that every OMQ from (EL, AQ) is in AC0, NL-complete, or PTIME-complete, a remarkable sparseness of complexities. We also establish elegant characterizations that separate the three classes of OMQs. In particular, we show that an OMQ Q is in NL if there is a bound k such that any minimal treeshaped ABox A whose root is an answer to the OMQ Q does not contain a full binary tree of depth k as a minor, and PTIMEhard otherwise. We additionally use a second, more operational characterization to determine the precise complexity of deciding whether a given OMQ is in AC0, NL-complete, or PTIME-complete, which turns out to be EXPTIME-complete.

And second, we put rewritability into linear Datalog onto the agenda of OMQ research. In fact, the equation "SQL = FO" often adopted in this area ignores the fact that SQL contains linear recursion from its version 3 published in 1999 on, which exceeds the expressive power of FO. We believe that, in the context of OMQs, linear Datalog is a natural abstraction of SQL that includes linear recursion, despite the fact that it does not contain full FO. Indeed, all OMQs from (EL, AQ) that are FO-rewritable are also rewritable into a union of conjunctive queries (UCQ) and thus into linear Datalog (and the same is true for much more expressive OMQ languages) [Bienvenu et al., 2014]. This shows that the expressive power of FO that lies outside of linear Datalog is not useful when using SQL as a target language for OMQ rewriting. We prove that rewritability into linear Datalog coincides with containment in NL. By what was said above, it is thus EXPTIME-complete to decide whether a given OMQ is rewritable. Moreover, we show how to construct linear Datalog rewritings when they exist and prove that there is no constant bound on the arity of IDB relations in linear Datalog rewritings.

Proof details are in the appendix, which is provided at http://www.cs.uni-bremen.de/tdki/research/papers.html.

## 2 Preliminaries

Let NC, NR, and NI be countably infinite sets of concept names, role names, and individual names. An EL-concept is built according to the syntax rule C, D ::= ⊤ | A | C ⊓ D | ∃r.C where A ranges over concept names and r over role names. An EL-TBox is a finite set of concept inclusions (CIs) of the form C ⊑ D, C and D EL-concepts. The size of T , denoted |T |, is the number of symbols needed to write all CIs of T , with each concept and role name counting as one symbol.

An ABox is a finite set of concept assertions A(a) and role assertions r(a, b) where A is a concept name, r a role name, and a, b individual names. We use Ind(A) to denote the set of individuals of the ABox A. A signature is a set of concept and role names. We often assume that the ABox is formulated in a prescribed signature, which we call the ABox signature. An ABox that only uses concept and role names from a signature Σ is called a Σ-ABox.

The semantics of DLs is given in terms of an interpretation I = (∆I , ·I ), where ∆I is a non-empty set (the domain) and ·I is the interpretation function, assigning to each A ∈ NC a set AI ⊆ ∆I and to each r ∈ NR a relation rI ⊆ ∆I ×∆I . The interpretation function is extended to compound concepts by setting ⊤I = ∆I , (C ⊓D)I = C I ∩DI , and (∃r.C)I = {d ∈ ∆ | ∃e ∈ ∆ : (d, e) ∈ rI }. An interpretation I satisfies a CI C ⊑ D if C I ⊆ DI , a concept assertion A(a) if a ∈ AI , and a role assertion r(a, b) if (a, b) ∈ rI . We say that I is a model of a TBox or an ABox if it satisfies all inclusions or assertions in it.

An atomic query (AQ) takes the form A(x), A a concept name. We write A, T |= A(a) if a ∈ Ind(A) and in every model I of A and T , we have a ∈ AI . An ontology-mediated query (OMQ) is a triple Q = (T ,Σ,A(x)) with T a TBox, Σ an ABox signature, and A(x) an AQ. We assume w.l.o.g. that A occurs in T . Let A be a Σ-ABox. We write A |= Q(a) and say that a is an answer to Q on A if A, T |= A(a). The evaluation problem for Q is to decide, given a Σ-ABox A and an a ∈ A, whether A |= Q(a). When we speak about the complexity of an OMQ Q, we generally mean its evaluation problem. It is thus understood what we mean when saying that Q is in PTIME or NL-hard. We use (EL, AQ) to denote set of all OMQs (T ,Σ,A(x)) where T is an EL-TBox. It is known that all OMQs in (EL, AQ) are in PTIME [Rosati, 2007; Krisnadhi and Lutz, 2007].

A Datalog rule ρ has the form S(x) ← R1(y1) ∧ · · · ∧ Rn(yn) where n > 0 and S, R1, . . . , Rn are relations of any arity and x, yi denote tuples of variables. We refer to S(x) as the head of ρ, and to R1(y1) ∧ · · · ∧ Rn(yn) as the body. Every variable that occurs in the head of a rule is required to also occur in its body. A Datalog program Π is a finite set of Datalog rules with a selected unary goal relation goal that does not occur in rule bodies. Relation symbols that occur in the head of at least one rule of Π are intensional (IDB) relations, and all remaining relation symbols in Π are extensional (EDB) relations. In our context, EDB relations must be unary or binary and are identified with concept names and role names. Note that, by definition, goal is an IDB relation. A Datalog program is linear if each rule body contains at most one IDB relation. The width of a Datalog program is the maximum arity of non-goal IDB relations used in it and its diameter is the maximum number of variables that occur in a rule in Π. For an ABox A that uses only EDB relations from Π and a ∈ Ind(A), we write A |= Π(a) if a is an answer to Π on A, defined in the usual way [Abiteboul et al., 1995].

A Datalog program Π over EDB signature Σ is a rewriting of an OMQ Q = (T ,Σ,A(x)) if for all Σ-ABoxes A and all a ∈ Ind(A), we have A |= Q(a) iff A |= Π(a). We say that Q is (linear) Datalog-rewritable if there is a (linear) Datalog program that is a rewriting of Q. It is well-known that, in EL, all OMQs are Datalog-rewritable. It follows from the results in this paper that there are simple OMQs Q = (T ,Σ,A(x)) that are not linear Datalog-rewritable, choose e.g. T = {∃r.A ⊓ ∃s.A ⊑ A} and Σ = {r, s, A}.

Throughout the paper, we generally and without further notice assume TBoxes to be in normal form, that is, to contain only concept inclusions of the form ∃r.A1 ⊑ A2, ⊤ ⊑ A1, A1 ⊓ A2 ⊑ A3, A1 ⊑ ∃r.A2 where all Ai are concept names. Every TBox T can be converted into a TBox T ′ in normal form in linear time [Baader et al., 2005], introducing fresh concept names; the resulting TBox T ′ is a conservative extension of T , that is, every model of T ′ is a model of T and, conversely, every model of T can be extended to a model of T ′ . Consequently, when T is replaced in an OMQ Q = (T ,Σ,A0(x)) with T ′ resulting in an OMQ Q′ , then Q and Q′ are equivalent in the sense that they give the same answers on all Σ-ABoxes. Thus, conversion of the TBox in an OMQ into normal form does not impact evaluation complexity nor rewritability into linear Datalog (or any other language).

We shall often deal with ABoxes that are tree-shaped. By a tree, we generally mean a directed (unlabelled) tree T = (V,E), defined in the usual way. Every ABox gives rise to a directed graph GA = (Ind(A), {(a, b) | r(a, b) ∈ A for some r}). We say that A is tree-shaped if GA is a tree and r(a, b), s(a, b) ∈ A implies r = s. The importance of tree-shaped ABoxes is due to the fact that OMQs from (EL, AQ) cannot distinguish between a Σ-ABox and its unraveling into a tree, see [Lutz and Wolter, 2012] or the appendix of this paper.

We introduce some further standard graph theoretic notions for ABoxes. A homomorphism from an ABox A1 to an ABox A2 is a total function h : Ind(A1) → Ind(A2) such that A(a) ∈ A1 implies A(h(a)) ∈ A2 and r(a, b) ∈ A1 implies r(h(a), h(b)) ∈ A2. We write A1 → A2 if there is a homomorphism from A1 to A2. A directed graph G = (V, E) is a minor of an ABox A if G is a minor of GA, that is, if G can be obtained from GA by deleting edges and vertices and by contracting edges. A path decomposition of a directed graph (V,E) is a sequence S1, . . . , Sn of subsets of V such that for every (a, b) ∈ E there is a set Si with a, b ∈ Si and Si ∩ Sk ⊆ Sj, for all i ≤ j ≤ k. A path decomposition is an (ℓ, k)-path decomposition if k = maxni=1 |Si| and ℓ = maxn-1 i=1 |Si ∩ Si+1|. The pathwidth of a directed graph (V, E) is the smallest k such that (V, E) has an (ℓ, k+1)-path decomposition for some ℓ ≥ 0. We identify the pathwidth of an ABox A with the pathwidth of GA.

## 3 NL, PTime, Linear Datalog Rewritability

We establish a dichotomy between PTIME and NL for evaluating queries from (EL, AQ), also showing that containment in NL coincides with rewritability into linear Datalog (unless NL = PTIME). The dichotomy is based on a characterization of containment in NL via a 'bounded amount of branching' in ABoxes whose root is an answer to the query. The linear Datalog programs constructed in the proofs are of unbounded width. We establish a hierarchy theorem which shows that this is unavoidable.

Let Q = (T ,Σ,A0(x)) be an OMQ. We say that Q is unboundedly branching if for every k ≥ 0, there is a treeshaped Σ-ABox A such that

1. A, T |= A0(a), a the root of A, and A is minimal with this property (w.r.t. set inclusion)
2. A has the full binary tree of depth k as a minor.

Otherwise, Q is boundedly branching. In the latter case, the branching limit of Q is the maximum integer k such that there is a tree-shaped Σ-ABox A that satisfies Conditions 1 and 2 above. The branching limit is 0 if there is no tree-shaped Σ-ABox A that satisfies Condition 1.

Example 1. (1) The OMQ Q1 = (T1, {A, r, s},A(x)) with T1 = {∃r.A ⊑ B1, ∃s.A ⊑ B2, B1 ⊓ B2 ⊑ A} is unboundedly branching as witnessed by the ABoxes A1, A2, . . . where Ai is a full binary tree of depth i, each left successor connected via the role name r, each right successor via the role name s, and with the concept name A asserted for each leaf.

(2) The OMQ Q2 = (T2, {A, r, s},B12(x)) with T2 = {∃r.A ⊑ B1, ∃s.A ⊑ B2, ∃s.B2 ⊑ B2,B1 ⊓ B2 ⊑ B12, ∃r.B12 ⊑ B1} is boundedly branching with branching limit one. In fact, every minimal tree-shaped Σ-ABox whose root is an answer to Q2 consists of a single r-path with an s-path starting at each non-leaf node and with A asserted for each leaf. Note that the number of individuals at which a branching occurs is unbounded in such ABoxes.

The following theorem sums up the results obtained in this section, except for the width hierarchy (Theorem 15).

Theorem 2. For every OMQ Q ∈ (EL, AQ), one of the following applies:

1. Q is PTIME-hard and not expressible in linear Datalog;

2. Q is rewritable into linear Datalog and thus in NL. Bounded branching of Q implies linear Datalog rewritability and delineates the two cases.

Note that Theorem 2 implies that any OMQ from (EL, AQ) is linear Datalog rewritable if and only if it is in NL (unless NL = PTIME). It is interesting to compare Theorem 2 with the result by [Afrati and Cosmadakis, 1989] that there are Datalog-queries that are not expressible as a linear Datalog program, but belong to NC 2 and are thus unlikely to be PTIME-hard.

## 3.1 Characterizations and PTime-Hardness

Theorem 2 provides a characterization of PTIME-hardness in terms of unbounded branching that is elegant, but does not lend itself to hardness proofs very well. For this reason, we establish a second characterization designed to enable a reduction from the PTIME-complete path systems accessibility (PSA) problem and show that both characterizations are equivalent. The new characterization will also be handy later on to decide the rewritability of OMQs into linear Datalog.

An instance of PSA takes the form G = (V, E, S, t) where V is a finite set of nodes, E is a ternary relation on V , S ⊆ V is a set of source nodes, and t ∈ V is a target node. G is a yes instance if t is accessible, where a node v ∈ V is accessible if v ∈ S or there are accessible nodes u, w with (u, w, v) ∈ E.

Before we can state the new characterization, we need some preliminaries. Let T be a TBox. A T -type is a set t of concept names from T that is closed under T -consequence, that is, if T |= ⊓t ⊑ A, then A ∈ t. For any ABox A and a ∈ Ind(A), we use tpA,T (a) to denote the set of concept names A from T such that A, T |= A(a), which is a T -type. If M is a set of concept names, then by M(a) we denote the ABox {A(a) | A ∈ M}. We also write A, T |= M(a), meaning that A, T |= A(a) for all A ∈ M. For every tree-shaped ABox A and a ∈ Ind(A), we use Aa to denote the sub-tree ABox of A that has a as the root. Moreover, we use Aa to denote A\Aa , that is, the ABox obtained fromA by removing all assertions that involve descendants of a (making a a leaf) and all assertions of the form A(a). We also combine these notations, writing for example Aabc for ((Aa )b)c.

Definition 3. An OMQ (T ,Σ,A0(x)) ∈ (EL,AQ) has the ability to simulate PSA if there are T -types t0, t1 and a treeshaped Σ-ABox A with root a and distinguished non-root individuals b, c, d where c and d are distinct incomparable descendants of b such that

<!-- formula-not-decoded -->

<!-- formula-not-decoded -->

Example 4. The OMQ Q1 from Example 1 has the ability to simulate PSA. Figure 1 shows a witnessing ABox A according to Definition 3 where t1 = {A, B1, B2} and t0 = {B2}.

PSA is PTIME-hard under FO-reductions [Immerman, 1999]. Using a reduction from this probem, we show that having the ability to simulate PSA is sufficient for PTIMEhardness under FO-reductions. In particular, we use the ABox A∧ from Definition 3 to implement an "and" gate where t0

Figure 1: Witness ABox for Example 4

![Image](https://ijcai.org/proceedings/2017/fulltext/164/image_000000_c0a9e5cf801373fcdd5015e2c7c9701b85019d533722fa607d6d625516ccb186.png)

and t1 represent the truth values zero and one to capture the behaviour of the ternary relation E in PSA.

Lemma 5. If Q ∈ (EL, AQ) has the ability to simulate PSA, then Q is PTIME-hard under FO-reductions.

To link Lemma 5 to Theorem 2, we next show that the ability to simulate PSA is equivalent to unbounded branching.

Proposition 6. Let Q ∈ (EL, AQ). Then Q has the ability to simulate PSA iff Q is unboundedly branching.

The "⇒" direction is proved by taking an ABox A that witnesses the ability to simulate PSA and then glueing together disjoint copies of A∧ to obtain tree-shaped ABoxes whose root is an answer to Q, which are minimal with this property, and that contain deeper and deeper full binary trees as a minor. The "⇐" direction is based on a combinatorial argument: if we take a minimal tree-shaped ABox that makes Q true at the root and contains a deep full binary tree as a minor, then it must contain an ABox that witnesses the ability to simulate PSA.

## 3.2 NL and Linear Datalog-Rewritability

We show that bounded branching characterizes containment in NL as well as linear Datalog rewritability, which therefore coincide (unless NL = PTIME). We also give a way to construct linear Datalog rewritings when they exist.

Proposition 7. Let Q ∈ (EL, AQ). Then Q is boundedly branching iff Q is rewritable into a linear Datalog program. Moreover, if the branching limit of Q is k, then there is a linear Datalog rewriting of width k + 1.

Direction "⇒". Let Q = (T ,Σ,A0(x)) be an OMQ from (EL,AQ). For each k &gt; 0, we construct a linear Datalog program ΠQ,k that is sound as a rewriting of Q and complete on ABoxes that do not have the full binary tree of depth k as a minor. The program ΠQ,k uses IDB relations of the form Pt1,...,tm where t1, . . . , tm, are T -types; the arity of this relation is m ≤ k. For any finite set S of concepts, we use clT (S) to denote the smallest (w.r.t. set inclusion) T -type t with T |= ⊓S ⊑ t. Let N be the set of all concept names from T . The program ΠQ,k consists of five types of rules:

Start rules: PclT (S)(x) ← S(x) for all S ⊆ N and where S(x) abbreviates ∧A∈S A(x);

Extension rules: Pt1,...,tm,clT (S)(x1, . . . , xm, y) ← Pt1,...,tm (x1, . . . , xm) ∧ S(y) for all S ⊆ N and T -types t1, . . . , tm;

Step rules: Pt1,...,tm-1,t(x1, . . . , xm-1, y) ← Pt1,...,tm (x1, . . . , xm) ∧ r(y, xm) ∧ S(y) for all S ⊆ N and T -types t1, . . . , tm where t = clT (S ∪ {∃r.A | A ∈ tm});

Consolidation rules: Pt1,...,tm-2,t(x1, . . . , xm-1) ← Pt1,...,tm (x1, . . . , xm-1, xm-1) for all S ⊆ N and T -types t1, . . . , tm, t where t = clT (tm-1 ∪ tm);

Goal rules: goal(x) ← Pt(x) for all T -types t with A0 ∈ t.

Example 8. We give a fragment of the programΠQ2,2 for the OMQQ2 from Example 1 that is equivalent to the full ΠQ2,2 and showcases the purpose of the different rules. For readability, we use representative concept names in the subscript of IDB relations instead of types:

<!-- formula-not-decoded -->

It can be verified that the program ΠQ,k is sound, that is, A |= ΠQ,k(a) implies A |= Q(a) for any Σ-ABox A. The following lemma states a form of completeness.

Lemma 9. If A is a tree-shaped ABox with root a0 that does not have the full binary tree of depth k as a minor andA, T |= A0(a0), then A |= ΠQ,k(a0).

Lemma 9 is proved by exhibiting a suitable strategy for applying the rules in ΠQ,k. Returning to the "⇒" direction of Proposition 7, we next show the following.

Lemma 10. If k-1 is the branching limit of Q, then ΠQ,k is a rewriting of Q.

The programs ΠQ,k allow us to construct a linear Datalog rewriting of an OMQ Q provided that we know an upper bound on its branching limit. The following lemma establishes such an upper bound (in case that Q is rewritable into linear Datalog at all).

Lemma 11. If Q = (T ,Σ,A0(x)) ∈ (EL, AQ) is boundedly branching, then its branching limit is at most 24|T |+1 .

In fact, Lemma 11 is a consequence of the proof of Proposition 6, given in the appendix. Lemma 11 almost yields decidability of linear Datalog rewritability: guess a tree-shaped Σ-ABox A and verify that it satisfies Conditions 1 and 2 from the definition of k-branching, where k is the bound from Lemma 11. For this to work, we would additionally have to bound the depth and degree of the tree-shaped ABoxes to be guessed. While this is not too difficult, we follow a different route (in Section 5) to obtain tight complexity bounds.

Direction "⇐". For d, k, n ≥ 0, let ℓkd (n) denote the maximum number of leaves in any tree that has degree d, depth n, and does not have as a minor the full binary tree of depth k + 1. The following lemma says that ℓkd (n) as a function of n grows like a polynomial of degree k.

Lemma 12. (d-1)k (n-k)k ≤ ℓkd (n) ≤ (k+1)(d-1)k nk for all d, k ≥ 0 and n ≥ 2k.

Let Π be a Datalog program over EDB signature Σ and IDB signature ΣI, and let A a Σ-ABox. It is standard to characterize answers to Π in terms of derivations that take the form of a labelled tree, see [Abiteboul et al., 1995] or the appendix. From each derivation D, one can read off an ABox AD in a standard way such that the properties summarized by the following lemma are satisfied.

Lemma 13. Let D be a derivation of Π(a) in A, Π of diameter d. Then

1. AD |= Π(a);
2. there is a homomorphismh fromAD toAwith h(a) = a;
3. AD has pathwidth at most d.

We are now ready to prove the desired result.

Lemma 14. If Q ∈ (EL, AQ) is unboundedly branching, then it is not rewritable into a linear Datalog program.

The proof, inspired by [Afrati and Cosmadakis, 1989], is by contradiction. Assume that Q ∈ (EL, AQ) is unboundedly branching, but rewritable into a linear Datalog program Π. We choose a minimal tree-shaped Σ-ABox A that contains a full binary tree of very large depth as a minor and such that A |= Q(a0), a0 the root of A. Consider the derivation D of Π(a0) in A and the associated ABox AD. By a sequence of manipulations, we identify a tree-shaped sub-ABox B ⊆ AD such that B has a very large number of leaves (a consequence of Point 2 of Lemma 13 and the fact that the homomorphism must be surjective due to the minimality of A and Point 1 of that lemma). By Lemma 12, it follows that B must contain a full binary tree of large depth as a minor and therefore must have high pathwidth, in contrast to Point 3 of Lemma 13.

## 3.3 Width Hierarchy

The linear Datalog rewritings constructed in the previous section are of unbounded width. We next show that this is unavoidable, in contrast to the fact that every OMQ from (EL,AQ) can be rewritten into a monadic Datalog program [Baader et al., 2017]. It strengthens a result by [Dalmau and Krokhin, 2008] who establish an analogous statement for constraint satisfaction problems (CSPs). However, while every OMQ from (EL, AQ) is equivalent to a CSP (up to complementation [Bienvenu et al., 2014]), the converse is false and indeed the CSPs used by Dalmau and Krokhin are not equivalent to an OMQ from (EL, AQ).

Theorem 15. For every ℓ &gt; 0, there is an OMQ from (EL, AQ) that is rewritable into linear Datalog, but not into a linear Datalog program of width ℓ.

To prove Theorem 15, we use the following queries: for all k ≥ 1, let Qk = (Tk, Σ, Ak(x)) where Σ = {r, s, t, u} and

<!-- formula-not-decoded -->

In the OMQ Qk, each concept name Ai, i ≤ k, represents the existence of a full binary tree of depth i, that is, if Ai is derived at the root of a tree-shaped Σ-ABox A, then A contains the full binary tree of depth i as a minor. Thus, deriving Qk at the root implies that A has the full binary tree of depth k as a minor.

Figure 2: An ABox of depth 4 whose root is an answer to Q2 and which is minimal with this property. It has 11 leaves, the largest number of leaves that a binary tree of depth 4 can have, unless it contains the full binary tree of depth 3 as a minor.

![Image](https://ijcai.org/proceedings/2017/fulltext/164/image_000001_b598c8c0f26aa5b62e7d29277a6eaca09191f3b3febd23307d639dfa9a0aff24.png)

Furthermore, for every n ≥ k there is minimal tree-shaped Σ-ABox A such that Qk is derived at the root, A is of depth n, and A has the maximum number of leaves that any tree of depth n without the full binary tree of depth k + 1 as a minor can have. For the case k = 2 and n = 4, such an ABox is shown in Figure 2. The concept inclusions ∃x.Bx,i ⊑ Bx,i in Tk ensure that Qk is closed under subdivisions of ABoxes, that is, if A is a Σ-ABox andA′ is obtained fromAby subdividing an edge into a path (using the same role name as the original edge), then A |= Qk(a) iff A′ |= Qk(a) for all a ∈ Ind(A).

Lemma 16. Every Qk is rewritable into linear Datalog.

We prove Lemma 16 by showing that each Qk is boundedly branching with branching limit k and using Proposition 7.

To show that linear Datalog rewritings of the defined family of OMQs require unbounded width, we first show that they require unbounded diameter and then proceed by showing that the width of rewritings cannot be significantly smaller than the required diameter. To make the latter step work, we actually show the former on an infinite family of classes of ABoxes of restricted shape. More precisely, for all i ≥ 0 we consider the class Ci of all forest-shaped Σ-ABoxes in which the distance between any two branching individuals exceeds i (where a forest is a disjoint union of trees and a branching individual is one that has at least two successors). Since the queries Qk are closed under the subdivision of ABoxes, each class Ci contains ABoxes whose root is an answer to the query. The proof of the following is similar to the proof of Lemma 14.

Lemma 17. For any i ≥ 0, Q2k+3 is not rewritable into a linear Datalog program of diameter k on the class of ABoxes Ci.

We are now ready to establish the hierarchy.

Proposition 18. Q8ℓ+13 is not rewritable into a linear Datalog program of width ℓ.

The proof of Proposition 18 is by contradiction. Assume that Q8ℓ+13 is rewritable into a linear Datalog program Π of width ℓ and let k be the diameter of Π. We show that, on the class of ABoxes Ck, there must then be a linear Datalog rewriting Π′ of Q8ℓ+13 of diameter 4ℓ+5, contradicting Lemma 17. In fact, Π′ can be obtained from Π by a sequence of manipulations: first rewrite the rules such that the restriction of rule bodies to EDB relations takes the form of a forest in which there is at most one branching node in every tree, then further rewrite to achieve that each such forest contains at most 2ℓ trees, and finally replace each rule with a set of rules of small diameter, slightly increasing the width.

## 4 AC0 vs. NL: Completing the Trichotomy

We say that an OMQ Q = (T ,Σ,A0(x)) has unbounded depth if for every k ≥ 0, there is a tree-shaped ABox A with depth at least k and root a such that A, T |= A0(a) and A is minimal with this property (regarding set inclusion). The following theorem summarizes the results in this section.

Theorem 19. For every OMQ Q ∈ (EL, AQ), one of the following applies:

1. Q is FO-rewritable and thus in AC0.
2. Q is not FO-rewritable and NL-hard.

Unbounded depth of Q implies NL-hardness and delineates the two cases.

The following characterization of FO-rewritability was established in [Bienvenu et al., 2013].

Theorem 20. Let Q ∈ (EL, AQ). Q is not FO-rewritable iff Q has unbounded depth.

To prove Theorem 19, it thus remains to show that unbounded depth implies NL-hardness. Similarly to the case of PTIME-hardness, the elegant condition of unbounded depth does not directly lend itself to hardness proofs, and we thus establish a second and equivalent characterization. Here, the second characterization is tailored towards NL-hardness proofs via reduction from reachability in directed graphs (REACH).

Definition 21. An OMQ (T , Σ, A0(x)) ∈ (EL, AQ) has the ability to simulate REACH iff there are T -types t0 ⊊ t1 and a tree-shaped ABox A with root a and distinguished non-root individuals b, c where c is a descendant of b such that

1. A, T |= A0(a),
2. t1 = tpA,T (b) = tpA,T (c),
3. tpAc∪t0(c),T (b) = t0, and
4. Ab ∪ t0(b), T ̸|= A0(a).

We define Afinish = Ab, Aedge = Abc , and Astart = Ac .

The three defined sub-ABoxes can be used in a reduction from REACH to Q. We now prove that unbounded depth implies NL-hardness, proceeding via the ability to simulate REACH. The following lemma is essentially implicit already in [Bienvenu et al., 2013].

Lemma 22. Let Q ∈ (EL, AQ). If Q has unbounded depth, then Q has the ability to simulate REACH.

The next lemma is proved similarly to Lemma 5.

Lemma 23. Let Q ∈ (EL, AQ). If Q has the ability to simulate REACH, then Q is NL-hard under FO-reductions.

We have completed the proof of Theorem 19, and thus also of the trichotomy.

## 5 Decidability and Complexity

We first show that an existing reduction in [Bienvenu et al., 2013] yields a variety of relevant hardness results, under various complexity-theoretic assumptions.

Theorem 24. The following properties of OMQs from (EL, AQ) are EXPTIME-hard: linear Datalog rewritability, containment in NL (unless NL = PTIME), NL-hardness (unless L = NL), and PTIME-hardness (unless L = PTIME).

For NL-hardness and PTIME-hardness, the complexitytheoretic assumptions can be dropped when hardness is defined under FO-reductions, as a consequence of the fact that Lemma 5 establishes hardness under such reductions.

Regarding upper bounds, we first recall the known result that it is in EXPTIME to decide whether an OMQ from (EL, AQ) is FO-rewritable [Bienvenu et al., 2013] and observe that, by Theorem 19, we also obtain an EXPTIME upper bound for NL-hardness. For linear Datalog rewritability, containment in NL, and PTIME-hardness, we use an approach based on (one-way) alternating parity automata on finite trees (APTAs). Because of space constraints, we can only give a brief sketch. By Theorem 2 and Proposition 6, it suffices to decide whether a given OMQ has the ability to simulate PSA, that is, whether there are T -types t0, t1 and a tree-shaped Σ-ABox A that satisfy the conditions from Definition 3. We iterate over all choices for t0, t1, building for each choice an APTA At0,t1 that accepts precisely the tree-shaped Σ-ABoxes satisfying the required conditions for the chosen t0, t1.

Theorem 25. It is in EXPTIME to decide whether a given OMQ from (EL, AQ) is rewritable into linear Datalog.

Interestingly, it is rather unclear how an EXPTIME upper bound would be established based on the characterization in terms of bounded branching. The following corollary sums up the results obtained in this section.

Corollary 26. For OMQs from (EL, AQ), all of the following problems are EXPTIME-complete (under the same complexity theoretic assumptions for the lower bounds as in Theorem 24): linear Datalog rewritability, containment in NL, NL-hardness, and PTIME-hardness.

Note that Theorem 25 and the results from Section 3.2 give an algorithm that provides a linear Datalog rewriting of a given OMQ if it exists and reports failure otherwise.

## 6 Conclusion

We plan to extend our analysis to (ELI, AQ) where ELI is the extension of EL with inverse roles. Then the overall picture changes because there are OMQs from (ELI, AQ) that express a form of undirected reachability and are L-complete. This also raises the question whether L-completeness coincides with rewritability into symmetric Datalog [Egri et al., 2007]. However, even lifting to (ELI, AQ) the results established in this paper such as the dichotomy between NL and PTIME is non-trivial. It would also be interesting to replace AQs with conjunctive queries. As illustrated by [Bienvenu et al., 2013] versus [Bienvenu et al., 2016], this makes the technical development more awkward since it requires to replace tree-shaped ABoxes with (somewhat contrived) ABoxes that are almost a tree. It might be more elegant to directly move to frontier-one tuple generating dependencies [Baget et al., 2009]. It would also be interesting to study the size of linear Datalog rewritings, to find ways to construct such rewritings that are efficiently executable, and to analyze empirically whether linear recusion is sufficiently well optimized in SQL database systems to support the rewritten queries.

## Acknowledgements

The authors were funded by ERC consolidator grant 647289.

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