3.1 Basic Definitions and Undecidability Results
Given countably infinite sets
\(\textsf {N}_{\textsf {C}}\) and
\(\textsf {N}_{\textsf {R}}\) of concept and role names,
\(\mathcal {ALC}\) concepts are built using the concept constructors top concept (
\(\top \)), bottom concept (
\(\bot \)), negation (
\(\lnot \)), conjunction (
\(C\sqcap D\)), disjunction (
\(C\sqcup D\)), existential restriction (
\(\exists r.C\)), and universal restriction (
\(\forall r.C\)). The semantics of the constructors is defined in the usual way (see, e.g., [
10,
12]). It assigns to every
\(\mathcal {ALC}\) concept
C a set
\(C^{\mathcal {I}} \subseteq \varDelta ^{\mathcal {I}} \), where
\(\varDelta ^{\mathcal {I}} \) is the interpretation domain of the given interpretation
\({\mathcal {I}}\). The set of
\(\mathcal {EL}\) concepts is obtained by restricting the available constructors to
\(\top \),
\(C\sqcap D\), and
\(\exists r.C\). As usual, a TBox is defined to be a finite set of
general concept inclusions (GCIs)
\(C\sqsubseteq D\), where
C,
D are concepts. The interpretation
\({\mathcal {I}}\) is a
model of such a TBox if
\(C^{\mathcal {I}} \subseteq D^{\mathcal {I}} \) holds for all GCIs
\(C\sqsubseteq D\) occurring in it. Given a concept description
C and a TBox
\({\mathcal {T}}\), we say that
C is
satisfiable w.r.t.
\({\mathcal {T}}\) if
\(C^{{\mathcal {I}}}\) is non-empty for some model
\({\mathcal {I}}\) of
\({\mathcal {T}}\). Concept satisfiability w.r.t. GCIs is ExpTime-complete in
\(\mathcal {ALC}\) [
71], but trivial in
\(\mathcal {EL}\) since
\(\mathcal {EL}\) concepts are always satisfiable. Given concept descriptions
C,
D and a TBox
\({\mathcal {T}}\), we say that
C is
subsumed by D w.r.t.
\({\mathcal {T}}\) (written
\(C\sqsubseteq _{\mathcal {T}} D\)) if
\(C^{\mathcal {I}} \subseteq D^{\mathcal {I}} \) holds for all models of
\({\mathcal {T}}\). Subsumption w.r.t. TBoxes in
\(\mathcal {ALC}\) is also ExpTime-complete since it interreducible with concept satisfiability, but tractable (i.e., decidable in polynomial time) in
\(\mathcal {EL}\) [
11,
32].
From an algebraic point of view, a concrete domain is a \(\tau \)-structure \(\mathfrak {D}\) with a relational signature \(\tau \) without constant symbols. To integrate such a structure into \(\mathcal {ALC}\) and \(\mathcal {EL}\), we complement concept and role names with a set of feature names \(\textsf {N}_{\textsf {F}}\), which provide the connection between the abstract domain \(\varDelta ^{\mathcal {I}} \) and the concrete domain D. A path is of the form \(r\,f\) or f where \(r\in \textsf {N}_{\textsf {R}}\) and \(f\in \textsf {N}_{\textsf {F}}\). In our example in the introduction, \(\textit{age}\) is both a feature name and a path of length 1, and \(\textit{child}\,\textit{age}\) is a path of length 2.
In contrast to previous works on concrete domains [
2,
65], we generally allow the use of the equality predicate in concrete domain restrictions, even if it is not explicitly contained in the signature of the concrete domain. This assumption will turn out to be useful later on, and it is basically without loss of generality since virtually all concrete domains considered in the literature can express equality in a way that does not impact on the complexity of reasoning. Our assumption that
\(\texttt {ff} \) is an atom implies that
\(\mathcal {EL} (\mathfrak {D})\) can express the bottom concept
\(\bot \) by the concrete domain restriction
\(\exists .\texttt {ff} \). A third difference is that, while features pointing into the concrete domain are functional, we do not allow the use of functional roles in paths. In [
2], only functional roles are allowed to occur in paths whereas in [
65] both functional and other roles can occur there. For
\(\mathcal {ALC}\), this does not really make a difference due to the availability of universal concrete domain restrictions. For
\(\mathcal {EL}\), the presence of functional roles would destroy tractability even without concrete domains [
11], and thus needs to be avoided anyway.
To define the semantics of concrete domain restrictions, we assume that an interpretation
\({\mathcal {I}}\) assigns functional binary relations
\(f^{{\mathcal {I}}}\subseteq \varDelta ^{{\mathcal {I}}}\times D\) to feature names
\(f\in \textsf {N}_{\textsf {F}}\), where
functional means that
\((a,d)\in f^{{\mathcal {I}}}\) and
\((a,d')\in f^{{\mathcal {I}}}\) imply
\(d=d'\). We extend the interpretation function to paths of the form
\(p = r\,f\) by setting
$$ (r\,f)^{{\mathcal {I}}} = \{(a,d)\in \varDelta ^{\mathcal {I}} \times D\mid \text{ there } \text{ is }\ b\in \varDelta ^{{\mathcal {I}}}\ \text{ such } \text{ that }\ (a,b)\in r^{{\mathcal {I}}}\ \text{ and }\ (b,d)\in f^{{\mathcal {I}}}\}. $$
The semantics of concrete domain restrictions is now defined as follows:
$$ \begin{array}{r@{}l} (\exists p_1,\dots ,p_k.\phi (x_1,\dots ,x_k))^{{\mathcal {I}}} = \{ a\in \varDelta ^{{\mathcal {I}}} \mid \text{ there } \text{ are }\ d_1,\ldots , d_k\in D\ \text{ such } \text{ that }\\ (a,d_i)\in p_i^{{\mathcal {I}}}\ \text{ for } \text{ all } i\in [k] \text{ and }\ \mathfrak {D} \models \phi (d_1,\dots ,d_k) \},\\ (\forall p_1,\dots ,p_k.\phi (x_1,\dots ,x_k))^{{\mathcal {I}}} = \{ a\in \varDelta ^{{\mathcal {I}}} \mid \text{ for } \text{ all }\ d_1,\ldots , d_k\in D\ \text{ such } \text{ that }\\ (a,d_i)\in p_i^{{\mathcal {I}}}\ \text{ for } \text{ all } i\in [k] \text{ we } \text{ have }\ \mathfrak {D} \models \phi (d_1,\dots ,d_k) \}. \end{array} $$
As already mentioned above, the concrete domain restriction
\(\exists .\texttt {ff} \) is unsatisfiable, and thus equivalent to
\(\bot \). The restriction
\(\exists f,f.(x_1=x_2)\) expresses that the value of the feature
f must be defined, without putting any constraint on this value.
Adding a concrete domain to a DL can easily lead to undecidability. Clearly, if the CSP of the concrete domain is undecidable, then this transfers to the DL it is integrated in. If the concrete domain is admissible, i.e., its CSP is decidable and its relations are closed under complements, then concept satisfiability without GCIs is decidable in a variant of
\(\mathcal {ALC} \) with concrete domains that uses functional roles in paths [
2]. But even for very simple concrete domains with decidable CSPs, the presence of GCIs may cause undecidability. For instance,
\(\mathcal {ALC} (\mathfrak {D})\) is undecidable already when
\(\mathfrak {D}\) is a structure over
\({\mathbb {N}}\) that has access to the unary predicate
\(=_{0}\), which is interpreted as the singleton set
\(\{0\}\), and the binary predicate
\(+_1\), which is interpreted as incrementation (i.e., it consists of the tuples
\((m,m + 1)\) for
\(m\in {\mathbb {N}}\)) [
12]. We can improve on this result by showing undecidability for even less expressive concrete domains without the predicate
\(=_{0}\).
This undecidability results also holds without our assumption that equality is always available, but the proof given in [
6] uses functional roles in paths. This proof can, however, easily be adapted to work also without functional roles. One simply must use additional universal quantification (i.e., value restrictions and universal concrete domain restrictions) to ensure that all the successors of an individual w.r.t. a role that was assumed to be functional in the original proof behave the same. More specifically, one must replace each occurrence of a concrete domain restriction of the form
\(\exists p_1,\dots , p_{k}.\phi \) for functional paths
\(p_i\) with the concept description
\(\exists p_1,\dots , p_{k}.\phi \sqcap \forall p_1,\dots , p_{k}.\phi \). Then the functionality restriction on these paths can be removed.
Even for \(\mathcal {EL}\), integrating a decidable concrete domain may cause undecidability if we allow for role paths of length 2. Proving this is, however, more challenging, not only due to the fact that not all Boolean operations are available, but also since the absence of functional roles cannot be compensated by the use of universal quantification. To illustrate the latter point, assume we have a concrete domain with binary predicates P and \(P'\) that are disjoint. If g is assumed to be a functional role, then the concept \( \exists f, g\,f .P(x_1,x_2)\sqcap \exists f, g\,f.P'(x_1,x_2) \) is unsatisfiable, but if g is just an arbitrary role, then it is satisfiable since a given individual belonging to the concept may have two different g-successors, one satisfying the P-constraint and the other satisfying the \(P'\)-constraint. However, conjoining this concept with the corresponding universal concrete domain restrictions \( \forall f, g\,f .P(x_1,x_2)\sqcap \forall f, g\,f.P'(x_1,x_2) \) yields an unsatisfiable concept again.
To show undecidability for a concrete domain extension of
\(\mathcal {EL}\) without functional roles, we consider the relational structure
\(\mathfrak {D}_{\mathrm {aff},{\mathbb {Q}} ^2}\) over
\({\mathbb {Q}}^2\), which has, for every affine transformation
\({\mathbb {Q}} ^2\rightarrow {\mathbb {Q}} ^2 : \bar{x} \mapsto A\bar{x}+\bar{b}\), the binary relation
\(R_{A,\bar{b}}:=\{(\bar{x},\bar{y})\in ({\mathbb {Q}}^2)^2 \mid \bar{y}= A\bar{x}+\bar{b} \}\) as a predicate. We will show in Corollary
8 that the CSP for this structure is decidable in polynomial time. Undecidability of subsumption w.r.t. TBoxes in
\(\mathcal {EL} (\mathfrak {D}_{\mathrm {aff},{\mathbb {Q}} ^2})\) can be shown by a reduction from
2-Dimensional Affine Reachability, which is undecidable by Corollary 4 in [
14]. For this problem, one is given vectors
\(\bar{v},\bar{w} \in {\mathbb {Q}}^2\) and a finite set
S of affine transformations from
\({\mathbb {Q}}^2\) to
\({\mathbb {Q}}^2\). The question is then whether
\(\bar{w}\) can be obtained from
\(\bar{v}\) by repeated application of transformations from
S.
Proof
We define the reduction of 2-dimensional Affine Reachability to subsumption w.r.t. general TBoxes in \(\mathcal {EL} (\mathfrak {D}_{\mathrm {aff},{\mathbb {Q}} ^2})\) as follows. For given vectors \(\bar{v},\bar{w} \in {\mathbb {Q}}^2\) and affine transformations \(S=\{\bar{x} \mapsto M_1\bar{x} + \bar{v}_1,\dots , \bar{x} \mapsto M_k\bar{x} + \bar{v}_k\}\), the TBox \({\mathcal {T}}\) contains, for every \(i\in [k]\), the GCI \(\top \sqsubseteq \exists f,g\,f.R_{M_i,\bar{v}_i}(x_1,x_2)\). Additionally, \({\mathcal {T}}\) contains the GCIs \(\exists g.L \sqsubseteq L\) and \(\exists f,f.R_{Z,\bar{w}}(x_1,x_2) \sqsubseteq L\), where L is a fresh concept name and Z is the \(2\times 2\) zero matrix. Note that \((\bar{x},\bar{x}) \in R_{Z,\bar{w}}\) iff \(\bar{x}=\bar{w}\). Each involved concept is either \(\top \), a concept name, or an existential (concrete domain) restriction, and thus definable in \(\mathcal {EL} (\mathfrak {D}_{\mathrm {aff},{\mathbb {Q}} ^2})\). We claim that \(\exists f,f.R_{Z,\bar{v}}(x_1,x_2)\) is subsumed by L w.r.t. \({\mathcal {T}}\) iff \(\bar{w}\) can be obtained from \(\bar{v}\) through an application of a composition of affine transformations from S.
“\(\Leftarrow \)”: Suppose that there exists such a composition and let \({\mathcal {I}}\) be a model of \({\mathcal {T}}\). Let a be an arbitrary element of \((\exists f,f.R_{Z,\bar{v}}(x_1,x_2))^{\mathcal {I}} \), i.e., satisfying \(f^{{\mathcal {I}}}(a)=\bar{v}\). Since \({\mathcal {T}}\) contains \(\top \sqsubseteq \exists f,gf.R_{M_i,\bar{v}_i}(x_1,x_2)\) for every \(i\in [k]\) and \(\bar{w}\) is reachable from \(\bar{v}\) through an application of a composition of affine transformations from S, there exists a role path \(a \rightarrow _{g^{{\mathcal {I}}}} \cdots \rightarrow _{g^{{\mathcal {I}}}} b\) to some element b with \(f^{{\mathcal {I}}}(b)=\bar{w}\). Since \({\mathcal {T}}\) contains the GCI \(\exists f,f.R_{Z,\bar{w}}(x_1,x_2) \sqsubseteq L\), we have \(b\in L^{{\mathcal {I}}}\). The GCI \(\exists g.L \sqsubseteq L\) then yields \(a\in L^{{\mathcal {I}}}\).
“\(\Rightarrow \)”: Suppose that \(\exists f,f.R_{Z,\bar{v}}(x_1,x_2)\) is subsumed by L w.r.t. \({\mathcal {T}}\). Consider the following interpretation \({\mathcal {I}}\). The domain of \({\mathcal {I}}\) is \({\mathbb {Q}}^2\). We define \(f^{{\mathcal {I}}}\) as the identity map on \({\mathbb {Q}}^2\) and set \( g^{{\mathcal {I}}}:=\{(\bar{x},\bar{y}) \in ({\mathbb {Q}}^2)^2 \mid \exists i\in [k] \text { such that } \bar{y}=M_i\bar{x}+\bar{v}_i \}\). Finally, we set \(L^{{\mathcal {I}}} :=\{\bar{w}\}\cup \{\bar{x} \in {\mathbb {Q}}^2 \mid \text {there exists a role path }\bar{x} \rightarrow _{g^{{\mathcal {I}}}} \cdots \rightarrow _{g^{{\mathcal {I}}}} \bar{w}\}\). It is easy to check that \({\mathcal {I}}\) is a model of \({\mathcal {T}}\). Since \(\bar{v}\in (\exists f,f.R_{Z,\bar{v}}(x_1,x_2))^{\mathcal {I}} \) and \(\exists f,f.R_{Z,\bar{v}}(x_1,x_2)\) is subsumed by L w.r.t. \({\mathcal {T}}\), we have \(\bar{v}\in L^{{\mathcal {I}}}\). The definition of \(L^{{\mathcal {I}}}\) thus implies that \(\bar{w}\) is reachable from \(\bar{v}\) through an application of a composition of affine transformations from S. \(\square \)
Note that the signature of \(\mathfrak {D}_{\mathrm {aff},{\mathbb {Q}} ^2}\) is infinite since there are infinitely many affine transformations on \({\mathbb {Q}} ^2\). One might think that this is important for our undecidability proof.
We can show, however, that this is not the case: a fixed finite set of affine transformations is sufficient (see the appendix for a proof).
3.2 Decidable and Tractable DLs with Concrete Domains
There are two strategies for regaining decidability of DLs with concrete domains in the presence of GCIs: syntactically restricting the interaction of the DL with the concrete domain or limiting the expressiveness of the concrete domain itself. Typically, the former is realized by restricting the length of paths in concrete domain restrictions to 1. We indicate this restriction by writing square brackets around the concrete domain instead of round brackets.
For
\(\mathcal {ALC}\), this restriction results in decidability [
45,
69] for concrete domains that are
admissible in the sense introduced in [
2], i.e., whose predicates are closed under negation and whose CSP is decidable. In the case of
\(\mathcal {EL}\), the expectations are a bit higher: the aim there is to regain tractability. To obtain tractability of
\(\mathcal {EL} [\mathfrak {D}]\), the notion of p-admissible concrete domains was introduced in [
11], and it was shown that subsumption in
\(\mathcal {EL} [\mathfrak {D}]\) is decidable in polynomial time iff
\(\mathfrak {D}\) is
p-admissible. Before defining this condition below, we introduce a condition, called
\(\omega \)-admissibility, which ensures decidability of
\(\mathcal {ALC} (\mathfrak {D})\) in the presence of GCIs and paths of length
\(> 1\).
3.2.1 \(\omega \)-Admissible Concrete Domains
The notion of
\(\omega \)-admissibility was introduced in [
65] to regain decidability of
\(\mathcal {ALC}\) with concrete domains in the presence of GCIs. Motivated by binary constraint calculi like Allen’s interval calculus [
1] and the region connection calculus [
70], only concrete domains where all predicates are binary were considered in [
65]. In [
6], the notion and the corresponding decidability result were generalized to concrete domains with predicates of arbitrary arity.
We say that the structure
\(\mathfrak {D}\) has
homomorphism \(\omega \)-
compactness if the following holds for every countable structure
\(\mathfrak {B}\):
\(\mathfrak {B}\rightarrow \mathfrak {D}\) iff
\(\mathfrak {A}\rightarrow \mathfrak {D}\) for every
\(\mathfrak {A}\in \mathrm {Age} (\mathfrak {B})\). In [
65], the inputs to this condition were not formally restricted to countable structures. However, it is clear that this is what the authors meant because (i) the structures produced by the original tableau algorithm that need to be tested for a homomorphism to the concrete domain are always countable, and (ii) the examples of
\(\omega \)-admissible concrete domains presented in [
65] are not homomorphic compact for arbitrarily large cardinalities. A relational
\(\tau \)-structure
\(\mathfrak {D}\) is
-
JE if, for every \(k\ge 1\), either \(\mathfrak {D}\) has no k-ary relations or \(\bigcup \big \{R^{\mathfrak {D}} \ \big | \ R\in \tau ,\ R^{\mathfrak {D}} \subseteq D^{k}\big \} = D^{k}\);
-
PD if \(R^{\mathfrak {D}}\cap {\tilde{R}}^{\mathfrak {D}}=\emptyset \) for all pairwise distinct \(R, {\tilde{R}}\in \tau \);
-
JD if equality
has a (quantifier and equality)-free definition in
\(\mathfrak {D}\).
Here JE stands for “jointly exhaustive,” PD for “pairwise disjoint,” and JD for “jointly diagonal.” In [
6], JD was defined in a more restricted way as
, which explains the name. The condition JD was not considered in [
65]. We include it here mainly because it makes the comparison with known notions from model theory easier. In the setting considered in the present paper, where concrete domain restrictions always have access to equality, JD is actually needed to ensure decidability. If the equality predicate is dropped from concrete domain restrictions, then the decidability results in [
7,
65] do not depend on JD. However, all examples of
\(\omega \)-admissible concrete domains presented in [
65] satisfy JD since equality is contained in the signature. In [
39],
k-ary structures, i.e., structures
\(\mathfrak {D}\) that have only
k-ary predicates, are considered that have the
k-ary equality relation
. For
\(k\ge 2\), such a structure satisfies JD in the sense introduced above, since binary equality
\(x=y\) can be defined as
.
A relational \(\tau \)-structure \(\mathfrak {D}\) is a patchwork if it is JDJEPD, and for all finite JEPD \(\tau \)-structures \(\mathfrak {A},\mathfrak {B}_{1}, \mathfrak {B}_{2}\) with \(e_{1}:\mathfrak {A}\hookrightarrow \mathfrak {B}_{1}\), \(e_{2}:\mathfrak {A}\hookrightarrow \mathfrak {B}_{2}\), \(\mathfrak {B}_{1}\rightarrow \mathfrak {D}\), and \(\mathfrak {B}_{2}\rightarrow \mathfrak {D}\), there exist \(f_{1}:\mathfrak {B}_{1}\rightarrow \mathfrak {D}\) and \(f_{2}:\mathfrak {B}_{2}\rightarrow \mathfrak {D}\) with \(f_{1}\circ e_{1}=f_{2}\circ e_{2}\).
The idea is now that one can use disjunctions of atomic formulas of the same arity within concrete domain restrictions. By \(\vee ^{\!+}\) we denote the set of all \(\tau \)-formulas of the form \(\phi _{1}(x_{1},\dots ,x_{k})\vee \cdots \vee \phi _{m}(x_{1},\dots ,x_{k})\) where each \(\phi _i\) is a k-ary atomic \(\tau \)-formula.
The following theorem is shown in [
6,
7] by extending the tableau-based decision procedure given in [
65] to our more general definition of
\(\omega \)-admissibility.
The main motivation for the definition of
\(\omega \)-admissible concrete domains in [
65] was that they can capture qualitative calculi of time and space. In particular, it was shown in [
65] that Allen’s interval logic [
1] as well as the region connection calculus RCC8 [
70] can be represented as
\(\omega \)-admissible concrete domains. To the best of our knowledge, no other
\(\omega \)-admissible concrete domains have been exhibited in the literature before our investigations in [
6], which we will describe in detail in the next section. Among other things, we prove that the structure
\(({\mathbb {Q}};<,=,>)\) is
\(\omega \)-admissible. The “discrete” version
\(({\mathbb {Z}};<,=,>)\), on the other hand, is not
\(\omega \)-admissible because it lacks homomorphism
\(\omega \)-compactness (see Example
1 below). By the results in [
60,
61],
\(({\mathbb {Z}};<,=,>)\) nevertheless yields a decidable concrete domain extension of
\(\mathcal {ALC} \), but proving this requires a more specialized approach than the tableau algorithm provided by the original paper of Lutz and Miličić [
65]. This shows that
\(\omega \)-admissibility is not necessary for decidable reasoning in
\(\mathcal {ALC} \) with concrete domains in the presence of GCIs.
3.2.2 p-Admissible Concrete Domains
The notion of p-admissibility was introduced in [
11] to capture precisely those structures
\(\mathfrak {D}\) for which subsumption in
\(\mathcal {EL} [\mathfrak {D}]\) is tractable. Clearly, this requires the CSP of
\(\mathfrak {D}\) to be decidable in polynomial time. However, this is not sufficient since even for a concrete domain
\(\mathfrak {D}\) with tractable CSP disjunction may be expressible in
\(\mathcal {EL} [\mathfrak {D}]\), which then leads to intractability [
11]. To avoid this source of intractability, the concrete domain must be convex. Unfortunately, the definition of convexity given in [
11] was ambiguous, and what is really needed in the setting considered in [
11] is what we call guarded convexity in [
8]. However, in the setting considered in the present paper, where equality is always available in concrete domain restrictions, we will see that convexity rather than guarded convexity is the adequate notion.
We say that a \(\tau \)-structure \(\mathfrak {D}\) is convex if the following holds: whenever a conjunction of atomic \(\tau \)-formulas implies a disjunction of atomic \(\tau \)-formulas in \(\mathfrak {D}\), then it already implies one of the disjuncts. Note that this definition does not say anything about which variables may occur in the left- and right-hand sides of such implications. Guarded convexity requires this condition to hold only for guarded implications, where all variables occurring on the right-hand side must also occur on the left-hand side.
To illustrate the difference between convexity and guarded convexity, let us consider the structure
\(\mathfrak {N} = ({\mathbb {N}};E,O)\) in which the unary predicates
E and
O are respectively interpreted as the even and odd natural numbers. This structure is not convex since
\(\forall x,y.(E(x)\Rightarrow E(y)\vee O(y))\) holds in
\(\mathfrak {N}\), but neither
\(\forall x,y.(E(x)\Rightarrow E(y))\) nor
\(\forall x,y.(E(x)\Rightarrow O(y))\) does. However, the first implication is not guarded, and it is easy to see that
\(\mathfrak {N}\) is in fact guarded convex. Note that, whereas
\(\forall x,y.\ (E(x)\Rightarrow E(y)\vee O(y))\) holds in
\(\mathfrak {N}\), the subsumption
\( \exists f.E(x_1) \sqsubseteq _\emptyset \exists g.E(x_1)\sqcup \exists g.O(x_1) \) does not hold in the extension of
\(\mathcal {EL} [\mathfrak {N}]\) with disjunction since the feature
g need not have a value. However, as we have pointed out above,
\(\exists g,g.(x_1= x_2)\) expresses that the value of
g must be defined. Thus,
\(\exists g,g.(x_1 = x_2) \sqsubseteq _\emptyset \exists g.E(x_1)\sqcup \exists g.O(x_1)\) is a valid subsumption in
\(\mathcal {EL} [\mathfrak {N}]\). This can be used to show that this DL is not tractable [
11], but only under the assumption that equality can be used in concrete domain restrictions. Consequently, in the setting of the present paper, convexity should be used in the definition of p-admissibility.
The main result of [
11] concerning concrete domains can now be stated as follows.
Note that the theorem above does not hold for the more expressive logic
\(\mathcal {EL}(\mathfrak {D})\) where paths of length 2 are allowed in concrete domain constructors. This is because we can show that the concrete domain
\(\mathfrak {D}_{\mathrm {aff},{\mathbb {Q}} ^2}\) from Proposition
2 is p-admissible (see Corollary
8). In Sect.
5, we provide an algebraic characterization of convexity. Regarding the tractability condition in the definition of p-admissibility, we have seen in Sect.
2 that it is closely related to the constraint satisfaction problem for
\(\mathfrak {D}\) and
\((\mathfrak {D}^{\lnot },\ne )\). Characterizing tractability of the CSP in a given structure is a very hard problem. Whereas the Feder-Vardi conjecture [
43] has recently been confirmed after 25 years of intensive research in the field by giving an algebraic criterion that can distinguish between
finite structures with tractable and with NP-complete CSP [
34,
72], finding comprehensive criteria that ensure tractability for the case of infinite structures is a wide open problem, though first results for special cases have been found (see, e.g., [
18,
25‐
28,
58,
67]).
3.2.3 \(\omega \)-Admissibility Versus p-Admissibility
From an application point of view it would be desirable to have concrete domains \(\mathfrak {D}\) that preserve tractability if used in \(\mathcal {EL}[\mathfrak {D}]\) and decidability if used in \(\mathcal {ALC}(\mathfrak {D})\). This would be the case for concrete domains that are both \(\omega \)-admissible and p-admissible. Unfortunately, for structures over a finite signature, JEPD (required for \(\omega \)-admissibility) and convexity (required for p-admissibility) do not go well together.
If
\(\tau \) contains a symbol
R that is interpreted as equality
on
D, then this proposition implies that
, which can only be the case if
\(|D|\le 1\). The proof of Proposition
3 makes use of our assumption that equality is always available when building formulas. But even without that assumption, concrete domains
\(\mathfrak {D}\) that are both p- and
\(\omega \)-admissible would have a rather restricted form. In that case, there always exists a finite partition
\(\{V_1,\dots , V_{m}\}\) of
D such that the only non-empty
k-ary relations of
\(\mathfrak {D}\) are of the form
\(V_{i_1}\times \cdots \times V_{i_k}\) for
\(i_1,\dots , i_{k}\in [m]\) [
5].
Finally, let us point out another notable difference between the two conditions, namely that p-admissibility permits infinite signatures whereas
\(\omega \)-admissibility does not. It turns out that finiteness of the signature is a necessary part of
\(\omega \)-admissibility to achieve decidability. If we allowed the signature of
\(\mathfrak {D}\) to be infinite, then we would have the following counterexample. Let
\(\mathfrak {D}\) be the structure over
\({\mathbb {Z}}\) with the relations
\(+_{k}^{\mathfrak {D}} = \{(x,y) \in {\mathbb {Z}}^2\mid y = x+k\}\) for every
\(k\in {\mathbb {Z}}\). It is easy to see that
\(\mathrm {CSP} (\mathfrak {D})\) can be solved in polynomial time and that
\(\mathfrak {D}\) has homomorphism
\(\omega \)-compactness. Moreover, one can show, using the results in Sect.
4 (Proposition
4 and Theorem
5), that
\(\mathfrak {D}\) is a patchwork. However, we have seen in Proposition
1 that concept satisfiability w.r.t. GCIs is undecidable already in
\(\mathcal {ALC} ({\mathbb {Z}};+_1)\).